Parallelizing Security Checks on Commodity Hardware
Edmund B. Nightingale∗†
Peter M. Chen∗
Dept. of EECS, University of Michigan∗
Speck1is a system that accelerates powerful security checks on
commodity hardware by executing them in parallel on multiple
cores. Speck provides an infrastructure that allows sequential in-
vocations of a particular security check to run in parallel without
sacrificing the safety of the system. Speck creates parallelism in
two ways. First, Speck decouples a security check from an appli-
cation by continuing the application, using speculative execution,
while the security check executes in parallel on another core. Sec-
ond, Speck creates parallelism between sequential invocations of a
security check by running later checks in parallel with earlier ones.
Speck provides a process-level replay system to deterministically
and efficiently synchronize state between a security check and the
original process. We use Speck to parallelize three security checks:
sensitive dataanalysis, on-access virusscanning, andtaint propaga-
tion. Running on a 4-core and an 8-core computer, Speck improves
performance 4x and 7.5x for the sensitive data analysis check, 3.3x
and 2.8x for the on-access virus scanning check, and 1.6x and 2x
for the taint propagation check.
Categories and Subject Descriptors
Reliability; D.4.6 [Operating Systems]: Security and Protection;
D.4.7 [Operating Systems]: Organization and Design; D.4.8 [Op-
erating Systems]: Performance
Performance, Reliability, Security
Operating Systems, Security, Performance, Parallel,
D.4.5 [Operating Systems]:
A run-time security check secures a system by instrumenting an
application to either detect an intrusion or prevent an attack from
succeeding. Security checks occur at a variety of granularities. For
instance, on-access virus scanners (Miretskiy et al. 2004) execute
during each read and write to disk, call-graph modeling (Wagner
and Dean 2001) executes during each system call, and security
checks such as taint analysis (Newsome and Song 2005), dynamic
data-flow analysis (Costa et al. 2005), and data-flow graph enforce-
ment (Castro et al. 2006) execute before a single instruction.
1Speck stands for Speculative Parallel Check.
Permission to make digital or hard copies of all or part of this work for personal or
classroom use is granted without fee provided that copies are not made or distributed
for profit or commercial advantage and that copies bear this notice and the full citation
on the first page. To copy otherwise, to republish, to post on servers or to redistribute
to lists, requires prior specific permission and/or a fee.
ASPLOS’08,March 1–5, 2008, Seattle, Washington, USA.
Copyright c ? 2008 ACM 978-1-59593-958-6/08/0003...$5.00
To detect an attack, a security check impedes the critical path of
an application by blocking it until the check completes. A security
check could be executed outside the critical path, but detecting an
attack after damage has been done is of little use – once an attack
succeeds, it is often very difficult to undo its effects.
Unfortunately, executing powerful security checks inthe critical
path drastically slows down performance. For example, a security
check such as taint analysis can slow down application execution
by an order of magnitude or more (Newsome and Song 2005).
Therefore, security researchers who wish to guard against intru-
sions are faced with a conundrum. Powerful security checks may
cause an application to execute too slowly, but weaker checks may
not detect an attack.
This paper describes a system, called Speck, that accelerates
powerful security checks on commodity hardware by executing
them in parallel on multiple cores. Speck provides an infrastructure
that allows sequential invocations of a particular security check to
run in parallel without sacrificing the safety of the system.
Speck creates parallelism in two ways. First, Speck decouples
a security check from an application by continuing the application,
using operating system support for speculative execution, while a
security check executes in parallel on another core. Second, Speck
takes advantage of the application executing ahead to start and
run later checks in parallel to earlier ones. Thus, Speck creates
parallelism not only between a process and a security check, but
also between the sequential invocations of a check.
Speck uses three techniques to safely defer and parallelize secu-
rity checks. First, Speck uses Speculator (Nightingale et al. 2005)
to execute a process speculatively and to track and propagate its
causal dependencies. Speculator allows a process to interact spec-
ulatively with the kernel, file system, and other processes. Should
a security check fail, Speculator rolls back each object and process
within the operating system that depends upon the compromised
Second, Speck buffers external output that depends on a specu-
lative process, such as output to the network or the screen, until the
security checks upon which the output depends have succeeded.
Finally, Speck provides a replay system that logs sources of
non-determinism as the speculative process executes. When the
security check executes in parallel, Speck replays each logged
source of non-determinism to ensure that the code executing in
parallel does not diverge from the speculative process.
We implemented Speck as a set of changes to the Linux 2.4 and
2.6 kernels, and we implemented parallel versions of three security
checks: sensitive data analysis, on-access virus scanning, and taint
propagation. Our results todate are promising. Running on a4-core
and an 8-core computer, Speck improves performance 4x and 7.5x
forthesensitivedataanalysischeck, 3.3xand2.8x fortheon-access
virus scanning check, and 1.6x and 2x for the taint propagation
. . .
. . .
. . .
The left-hand side shows a process executing with a security check
interrupting execution to check for an intrusion. The right-hand side
of the figure shows how Speck parallelizes security checks. On
CPU0, the original process runs, using speculative execution, without
a security check. Speck forks copies of the process, and uses replay
to aggregate and execute checks in parallel on multiple cores.
Figure 1. Parallelizing a sequential security check
Speck’s goal is to provide security equivalent to that provided by
security checks executing sequentially within the critical path of
an application, while achieving better performance through paral-
lelism. The security provided by a parallel check using Speck is
equivalent to that provided by a sequential version if the output
of the parallel version could have been produced by the sequential
The left-hand side of Figure 1 shows a program (denoted by the
numbered boxes) and a security check (striped boxes) executing
sequentially. The right-hand side of the figure shows how Speck
parallelizes security checks on multiple cores. The process running
on CPU0executes without a security check using operating sys-
tem support for speculative execution. We refer to it as the unin-
strumented process. The uninstrumented process invokes security
checks that run in parallel on other cores. Speck replays the unin-
strumented process’s execution (shaded boxes) on other cores to
generate the state required for each security check to run. We refer
to the processes that run security checks in parallel as instrumented
The state used by a parallel check must be identical to the state
it would have used if it was executing sequentially. When it is time
to begin a security check, Speck uses fork as a general method to
copy (i.e., copy-on-write) the state of the uninstrumented process
and ship it to another core. Using a software mechanism such as
fork allows Speck to run on commodity hardware. Unfortunately,
two non-trivial costs slow the start of a security check when using
fork: first, the page table of the application must be copied and
the pages must be protected. Second, during execution, both the
uninstrumented and cloned process will take copy-on-write page
Speck amortizes the cost of fork by aggregating multiple se-
curity checks into intervals of time called epochs. Increasing the
epoch length reduces the frequency of fork and thus reduces the
overhead needed to parallelize the check. However, increasing the
epoch length forces the uninstrumented process to execute further
ahead to begin the next epoch. For many applications, a longer
epoch increases the chances that the uninstrumented process will
execute a system call. Allowing the uninstrumented process to ex-
ecute system calls led to a three challenges that shaped the design
First, some system calls (e.g., write) allow a process to af-
fect other processes or kernel state. It is safe to allow a speculative
process to affect other processes and kernel objects as long as the
affected state is rolled back when a security check fails. We use
Speculator (Nightingale et al. 2005) to track and undo all changes
a speculative process makes to other processes and kernel objects.
Speculator tracks and propagates the causal dependencies of spec-
ulative processes as they execute. For example, if a speculative pro-
cess writes to a file, that file becomes speculative and inherits the
dependencies of the process. Each object and process within the
kernel has an undo log associated with it. If a security check fails,
Speculator uses the undo log to roll back each object and process
that depends on the failed security check. Thus, a compromised
program cannot permanently damage the system.
Second, other system calls allow a process to execute an output
commit (output that cannot be undone, e.g., writing to the screen).
When a process attempts an action that cannot be undone, the pro-
cess must block until all security checks it depends on are resolved.
Some output commits can be deferred, rather than blocked, which
improves performance by allowing the speculative process to con-
tinue executing further ahead. For example, Speculator buffers out-
put tothe screen and the network until the output no longer depends
upon speculative state. When Speculator cannot propagate a de-
pendency or buffer output, the process is blocked until all security
checks up to that point have passed.
Finally, a set of system calls exist that affect the state of the
calling process (e.g., read). Executing one of these system calls in-
troduces non-determinism into the execution of the uninstrumented
process. The non-determinism could cause the instrumented clone
to diverge from the path taken by the uninstrumented process. This
could cause the check to miss an attack experienced by the unin-
strumented process. Speck provides a transparent kernel-level re-
play system, which logs sources of non-determinism while the
uninstrumented process executes ahead. Speck replays each source
of non-determinism, and therefore the instrumented clone executes
the same code path as its uninstrumented counterpart.
Deterministic replay enables an instrumented clone to follow
the same progression of states as the uninstrumented process. Re-
play provides state to a subsequent security check equivalent to the
state provided if the uninstrumented process calls fork, but at a
much lower overhead. Thus, fork is the necessary mechanism to
create parallelism, and replay is the necessary mechanism to amor-
tize the cost of fork through the aggregation of security checks.
Fork and replay can efficiently provide the entire state of the
uninstrumented process to an instrumented clone, and this allows
one to run arbitrary security checks. However, some expensive
security checks run infrequently or require little state to execute.
In these cases, it may be faster to block the uninstrumented process
and send the required state directly to a separate security check
program, rather than use fork and replay to generate the state in an
instrumented clone. For these types of checks, Speck can disable
fork and replay and instead send only the subset of state required
for the check (see Section 4.3 for an example).
Two mechanisms ensure that Speck provides security equivalent to
a sequential security check. First, due to Speck’s replay system, the
instrumented clones will execute the same sequential checks that
wouldhavebeen executedsequentially. Second, due toSpeculator’s
causal dependency tracking and rollback, any actions a speculative
process takes before the instrumented clones complete the prior
checks are rolled back should a prior check fail.
Thus, Speck’s security guarantee assumes that an attacker can-
not compromise the replay system provided by Speck or the causal
dependency tracking and rollback system provided by Speculator.
Because Speck and Speculator operate within the kernel, a spec-
ulative process (which may be compromised but not yet checked)
must not be allowed to gain arbitrary control of kernel state. Specu-
lator’s design philosophy helps prevent speculative processes from
damaging thekernel. Bydefault, Speculator blocksspeculative pro-
cesses when they make system calls; only specific system calls to
specific devices are allowed to proceed. For example, Speculator
allows write system calls only to specific devices. Thus, while
the uninstrumented process may execute speculatively beyond the
point of the check, it should be unable to execute actions that com-
promise replay or Speculator (e.g., by writing to /dev/mem).
We do not prevent attacks that would compromise the kernel
while running a sequential version of the security check, as such
attacks would compromise the system with or without Speck.
We have implemented Speck within the Linux 2.4 and 2.6 kernels.
We provide a description of the implementation of epochs, the
replay system, and a brief overview of Speculator.
Speck must choose the epoch length to balance the amount of work
completed by the security check and the cost of beginning a new
epoch. If the epoch length is too short, the cost of fork dominates
the benefits derived from dividing the work of the security check
among multiple cores. Alternatively, if the epoch length is too long,
there are fewer opportunities to create parallelism in the system,
since the uninstrumented process may execute calls (e.g., select)
that forces Speck to wait as well.
We implemented Speck to balance the cost of starting an epoch
with the work accomplished during an epoch by using a time
threshold. At the end of each system call, Speck checks whether
the amount of time the uninstrumented program has executed is
greater than a threshold epoch length; if so, Speck terminates the
current epoch andforks anew instrumented clone. Aninstrumented
clone exits when it executes the last system call in its epoch. In our
experiments, we have found that an epoch length of 25–50ms is
short enough to create sufficient parallelism, but long enough to
provide a significant performance improvement.
Creating an epoch
Speck uses deterministic replay to ensure that the instrumented
and uninstrumented clones execute identical code paths, absent the
additional security checks in the instrumented version. If the two
code paths diverged, the security checks would not be valid.
When a program executes on a single core, there are only
two sources of non-determinism: system calls and signal deliv-
ery. Scheduling and hardware interrupts are transparent to appli-
cation execution. After a scheduling or hardware interrupt occurs,
an application resumes execution at the same point that it was in-
terrupted. Thus, operating system virtualization, which gives the
application an illusion of executing alone on a virtual processor,
masks this source of non-determinism from the application as long
Ensuring deterministic execution
asall forms of inter-process communication arelogged. Speckadds
support for deterministic replay to the Linux kernel by logging sys-
tem call results and signal delivery (Bressoud and Schneider 1995;
Dunlap et al. 2002). Our current implementation of deterministic
replay does not support shared memory IPC for multi-threaded
programs running on multiple cores. Speck could be modified to
support shared memory by adding support for deterministic, multi-
processor replay. Without specialized hardware support (Xu et al.
2003), Speck would need to fault on each shared memory access
and record the location and effect of the fault. This technique could
be prohibitively slow for workloads with many shared memory
Speck associates a save and a restore function with each system
call. When an uninstrumented program completes a system call,
Speck executes the associated save function to log the system call
results. When an instrumented clone makes a system call, Speck
redirectsit toasecondary replay systemcall table. Eachentry inthe
replay system call table points to an associated restore function
for that call, which returns the values logged by the uninstrumented
version of the program. For some system calls, such as nanosleep
andrecv, replayingtheloggedvalueintheinstrumented clonemay
take much less time than the original system call.
When an uninstrumented process begins a new epoch, Speck
creates a replay queue data structure for that epoch. Each replay
queue is shared between the uninstrumented and instrumented pro-
cesses for a given epoch. The two processes have a producer-
consumer relationship. The uninstrumented process adds system
call results to the FIFO queue, and its instrumented counterpart
pulls those results from the queue as it executes.
Unfortunately, this simple save and replay strategy is not
sufficient for all system calls. System calls such as read and
gettimeofday require special strategies because they return data
by reference rather than by value, i.e., they copy data from the
kernel into the user-level program’s address space. To replay such
calls, Speck saves the value of all user-level memory that was
modified by the execution of the system call. For example, Speck
saves the contents of the buffer modified by read. When the in-
strumented process calls the restore function for such system calls,
Speck copies the saved values from the replay queue into the appli-
Currently, Speck does not support processes that execute nonde-
terministic functions that bypass the operating system (e.g., RAW
sockets or the processor timestamp instruction rdtsc). Such ac-
tions could be supported either by logging them via binary instru-
mentation or by making the operations privileged and forcing the
process to enter the kernel.
System call replay
Certain functions such as mmap and brk modify the layout of an ap-
plication’s address space. Speck must re-execute these system calls
when they are called by an instrumented clone. However, Speck
ensures that the modification to the instrumented version of the ad-
dress space is identical to the modification made to the uninstru-
mented version. For example, the mmap system call is passed an ad-
dress as a suggested starting point to map the object into its address
space. If this parameter is NULL, then the operating system has the
freedom to choose any unused point in the program’s address space
to map memory. Naively replaying this call might cause the oper-
ating system to map memory to two different locations in different
copies of program. To prevent this problem, Speck saves the ad-
dress returned by mmap instead of the NULL value of the parameter.
When the instrumented clone later executes the same mmap system
call, Speck passes the saved address into the sys mmap function.
System call re-execution
This ensures that the memory is mapped to identical locations in
Linux delivers signals at three different points in process execution.
First,if aprocess issleeping ina systemcall such asselect, Linux
interrupts the sleep and delivers the signal when the process exits
thekernel. Second, Linux deliversthe signal when aprocess returns
from a system call. Finally, Linux delivers a signal when a process
returns from a hardware interrupt.
Toensuredeterministicexecution, Speckmust replaythereceipt
of a signal at the same point in the instrumented clone as it was re-
ceived by the uninstrumented process. Delivering a signal within
calls such as select or on the return from a system call is han-
dled by checking during system call replay whether a signal must
be delivered. The delivered signals are saved in the replay queue
during the execution of the uninstrumented process, and identical
signals are delivered when the instrumented clone returns from the
same system calls. Since the instrumented clone returns immedi-
ately from functions such as select rather than sleep, Speck never
needs to wake up instrumented processes in the kernel to deliver
Replaying signal delivery after a hardware interrupt is more dif-
ficult. To ensure correctness, a signal must be delivered at exactly
the same point of execution (e.g., after some identical number of
instructions have been executed). Our current implementation does
not yet support this functionality, although we could use the same
techniques employed by ReVirt (Dunlap et al. 2002) and Hyper-
visor (Bressoud and Schneider 1995) to deterministically deliver
such signals. Currently, Speck simply restricts signal delivery to
occur only on exit from a system call or after interrupting a system
call such as select. In practice, this behavior has been sufficient
to run our benchmarks.
We use Speculator (Nightingale et al. 2005) to provide support
for speculative execution within the Linux kernel. Speculator can
checkpoint the state of a process when it is executing within a
system call and execute it speculatively.
Speculator ensures that speculative state is never visible to an
external observer. If a speculative process executes a system call
that would normally externalize output, Speculator buffers its out-
put until the outcome of the speculation is decided. If a speculative
process performs a system call that Speculator is unable to han-
dle by either transferring causal dependencies or buffering output,
Speculator blocks it until it becomes non-speculative.
Speck uses Speculator to isolate an untrusted application until
a security check determines the application has not been compro-
mised. Speckassociates each epoch withaspeculation. Thesuccess
or failure of the security checks associated with an epoch deter-
mines whether the execution of code during that epoch was safe or
whether an intrusion occurred. If execution was safe, Speck com-
mits the speculation, which allows Speculator to release all output
generated by that epoch. If an intrusion occurred, Speck fails the
speculation, which causes Speculator to begin rollback. Each ob-
ject and process dependent upon the compromised application is
rolled back to a point before the intrusion occurred. We originally
considered rolling back the compromised application as well, but
realized that a sequential security check would likely terminate the
application after detecting an intrusion. Therefore, Speculator ter-
minates a compromised application. Since the application is never
rolled back, Speculator does not checkpoint its state at the begin-
ning of each epoch.
Currently, there is a slight difference in the output seen by an
external observer when a security check fails using Speck and the
OS support for speculative execution
output that would have been generated by a sequential security
check. Since Speck operates on the granularity of an epoch, any
output that occurred after the last epoch began but before the
instruction that led to thefailed security check would not be visible.
We plan to address this by simply replaying the uninstrumented
process up to the point where the security checks failed.
When Speck begins a new instrumented clone, the structure of
the clone depends upon the type of security check. Speck can
use any particular method of inserting checks. Two of the case
studies (4.2 and 4.4) in this paper use Pin (Luk et al. 2005) to
dynamically instrument the code with the necessary checks. Pin
allows application programmers to create Pintools, which contain
arbitrary code and rules about when that code should be inserted
into a program binary. Pin then uses the Pintool in tandem with
dynamic compilation to generate new program code.
Pintools are transparent to the execution of the application in-
strumented with Pin. An application instrumented with Pin ob-
serves the same addresses and registers for program code and data
as it would were it running without Pin. Speck uses this property
to ensure correctness: the execution of the instrumented clone does
not diverge from the execution of the uninstrumented process be-
cause the program cannot observe that it is being instrumented. The
authors of Pin have shown that Pin is faster than other systems such
as Valgrind (Nethercote and Seward 2007) and DynamoRIO (Sul-
livan et al. 2003). Since Pin does not interrupt program execution
unless a rule within a Pintool instructs it to do so, the overhead of
Pin without a Pintool is negligible.
Pin: A dynamic binary rewriter
In this section we discuss the properties of a security check that
determine both the difficulty in parallelization and the potential
for performance improvement through increased parallelism. We
discuss three different classes of security checks that demonstrate
tradeoffs in the difficulty and overhead of parallelization and then
describe the design and implementation of a check from each class.
Parallelizing security checks
When security checks are parallelized to run using Speck, checks
that would otherwise run in sequential order execute concurrently.
easier to parallelize) than others. This section describes the prop-
erties that determine the likelihood of success when parallelizing a
check using Speck.
Speck uses the uninstrumented process to run ahead and start
later security checks in parallel with earlier ones. Therefore, each
security check depends upon the state received from the uninstru-
mented process. This dependency impacts the amount of paral-
lelismSpeck canachieve intwoways. First,theamount of work ex-
ecuted by the uninstrumented process between checks determines
how quickly future epochs will begin. Second, the expense of the
security check determines how many checks might execute in par-
allel at any one time. Expensive checks create more opportunity for
parallelism and therefore offer the opportunity for greater relative
The benefit of parallelization also depends upon whether a later
check depends upon the result of an earlier check. If a later check
depends heavily on an earlier check running in parallel, it may
be difficult to parallelize the check in such a way that the later
check can make forward progress. If a later check is completely
independent of anearliercheck, thecheck isvery easytoparallelize
as the only dependency is the state provided by Speck from the
Choosing security checks
Incidentally, there is a third type of dependency relationship to
consider when parallelizing a security check. Some checks may
have very few dependencies within a single invocation of a check.
Such a check may be embarrassingly parallel, and could be par-
allelized within a single invocation, rather than parallelized across
multiple invocations as is done by Speck. Speck does provide one
advantage, which is that the programmer does not need to reason
about how to parallelize the check; parallelization is achieved au-
While signs of many security problems appear in the address space
of a process, these signs may appear in memory only for a brief
period of time. For example, a virus may decrypt itself, damage the
system, then erase itself (Szor 2005); sensitive, personal data may
inadvertently leak into an untrusted process (Chow et al. 2005); or
a malicious insider may release classified data by encrypting it and
attaching it to an e-mail. One could detect such transient problems
by checking memory at each store instruction, but the overhead of
such frequent checks would be prohibitive.
Speck can reduce the overhead of these checks by running
them in parallel. Parallelizing this type of check with Speck is
straightforward because the checks are independent of one another.
We implemented one such security check, which looks for in-
advertent leaks of sensitive data. The security check carries the 16-
byte hash of a sensitive piece of data. At each memory store, it cal-
culates the hash at all 16-byte windows around the address being
stored. We implemented the security check as a Pintool, which can
berun sequentially orinparallel by using Speck. Weused fork and
replay to efficiently synchronize state between the security check
and the uninstrumented process.
Process memory analysis
There have been many different proposed security checks within
the research community that analyze program behavior using sys-
tem calls. Examples include Wagner and Dean’s (2001) call graph
modeling, Provos’ (2003) systrace, Ko and Redmond’s (2002) non-
interference check, and on-access virus scanning (Miretskiy et al.
2004). These checks can dramatically slowdown performance,
which limits their usefulness in production environments. Wagner
and Dean cite multiple orders-of-magnitude slowdown, and Provos
notes a 4x slowdown in the worst case.
System call analysis exhibits two traits that make it promising
for use with Speck. First, each check does not depend upon the re-
sult of prior checks. Therefore, later checks can be run in parallel
with earlier checks without modification. Second, system call anal-
ysis requires little state to execute – often just the list of prior calls
executed and/or the parameters passed by the application. This trait
means that such checks do not require fork and replay; Speck can
ship the state required at each system call from the uninstrumented
process to an instance of the check executing in parallel.
We have implemented an on-access virus scanner using scan-
ning libraries provided by ClamAV (2007), a popular anti-virus
toolkit for Linux. Our on-access scanner intercepts file system re-
lated system calls, blocks the calling process, and passes the name
of the file down to a waiting scanning daemon. Once the file has
been checked, the process is allowed to continue executing.
We parallelize the check by creating a pool of waiting virus
scanning daemons. Since only the name of the file is required to
execute the scan, fork and replay are not used. Instead, our on-
access scanner sends the name of the file to the next available
scanner, and allowsthe process to continue executing speculatively.
System call analysis
Taint analysis traces the flow of data from untrusted sources (e.g.,
a socket) through an application’s address space and detects if a
critical instruction executes with tainted (i.e., untrusted) data. For
example, if the data on the top of the stack used by RET is tainted,
a stack smash attack has occurred.
As a process executes, the taint analysis check updates a map
representing which addresses and registers are tainted within the
process’s address space. Each check depends upon an updated ver-
sion of the map. Therefore, each check depends upon all prior
checks. These inter-check dependencies make taint analysis dif-
ficult to parallelize. Running multiple checks in parallel will not
work, since a later check executing in parallel to an earlier check
willnot yet havealltheinformationnecessary todetermine whether
an attack has occurred.
Because taint analysis is difficult to parallelize, we use it as a
challenging case study to parallelize with Speck. Our strategy is
to divide work into parallel and sequential phases of execution.
The parallel phase executes within instrumented clones, and the
sequential phase processes the parallel pieces to determine whether
an attack has occurred. Our goal is to push as much work into
the parallel phase as possible, while minimizing the work done
sequentially. The rest of this section gives a broad overview of
instruction-level taint analysis on x86 processors.
Prior taint analysis systems (Costa et al. 2005; Newsome and
Song 2005; Qin et al. 2006) use binary rewriters to instrument
a target application. A memory or register location is marked as
tainted if it causally depends on data from an external source (e.g.,
the disk or a network socket). For example, if an application reads
data from the network, a taint tracker would mark the memory
locations into which the data is copied as tainted. Data movement
instructions (e.g., MOV, ADD, and XCHG) propagate taint from one
location to another. Such instructions may also clear an address of
taint either by overwriting it with untainted data or by executing
an explicit clear instruction (e.g., XOR eax, eax). Taint trackers
must set input, propagation, and assert policies to determine when
data becomes tainted, which instructions propagate (or clear) taint,
and when to check for an intrusion, respectively.
An input policy determines when data from outside the address
space taints data within it. The input policy for our algorithm
taints data from any source external to the process’s address space,
excepting shared libraries loaded by the dynamic loader. Setting an
input policy can be reduced to a question of trust: if the loader and
shared libraries are trusted, then loading a shared library does not
taint the address space.
A propagation policy determines which instructions propagate
taint from one location to another. For example, taint analysis
functions typically instrument each data movement instruction to
propagate taint. Taint analysis tools may also instrument control
instructions such as JMP; e.g., a conditional jump based on a tainted
value could beconsidered totaintall datavalues modified later.Our
propagation policy instruments all data movement (MOV, MOVS,
PUSH, POP, ADD, CMOV, etc.) instructions withinthe application,
including F86 and SSE instructions, to propagate taint from data
sources to data sinks.
There is a tradeoff in determining a propagation policy between
the coverage provided and the accuracy of the algorithm. If taint
is propagated along all channels of information flow, taint analysis
will generate many false positives. However, if too many channels
of information flow are ignored, then attacks may go undetected.
Since Speck is independent of the specific taint analysis check, we
chose policies from prior work (Costa et al. 2005; Newsome and
Finally, an assert policy determines when to check whether
an attack has occurred. Assert policies share the same cover-
. . .
. . .
Log Preparation Log Processing
This figure shows the taint analysis check parallelized using Speck.
Each check depends upon the results of all prior checks. Therefore,
we invented a new parallel algorithm that divides the work into
parallel and sequential phases.
Figure 2. Parallelizing taint analysis using Speck
age/accuracy tradeoff that is inherent in propagation policies. We
implemented two assert policies – checking for stack smashing
attacks and function pointer attacks.
Since the sequential and parallel taint trackers required two dif-
ferent implementations, we divide the discussion into two separate
sections. Section 5 describes our sequential taint analysis check,
and Section 6 describes our parallel taint analysis check.
We built a sequential taint analysis check as a baseline comparison
by following the descriptions written by Costa et al. (2005) and
Newsome and Song (2005).
Our sequential taint tracker uses a page table data structure to
trackthetainted memoryaddresses of theinstrumented application.
The top level of the page table is an array of pointers. Each entry
points to a single page of memory in the instrumented application’s
address space. Each 4KB page is a character array, such that each
byte in the array represents a single byte of application memory.
We use a byte rather a bit array because our initial results showed
that bit operations substantially slowed down the taint tracker. To
save space, initially each entry of the page table points to the same
zero page, which is a 4 KB zero-filled array (indicating that the
corresponding memory addresses are untainted). The taint values
of registers are stored in another small array.
We use Pin to instrument each instruction that is covered by
an input, propagation, or output policy. Since Pin in-lines Pintool
code that is constructed without conditionals, the instrumentation
that comprises the sequential taint tracker is carefully constructed
to avoid conditional instructions.
Although our focus is on the benefits of parallelization, we
strove to efficiently implement Newsome and Song’s sequential
taint tracker using Pin. Our results show that our implementation
runs approximately 18 times slower than native speed; in compari-
son, Newsome and Song (2005) reported that their implementation
ran 24-37 times slower on an identical benchmark. Based on this,
we believe that our sequential implementation is reasonably effi-
Sequential taint analysis implementation
The need to sequentially update a map of tainted addresses and
registers as aprocess executes poses a challenge for Speck; running
instrumented clones in parallel would result in incorrect results,
since each epoch would have an incomplete map of taintedmemory
locations. We realized that the taint analysis check has two parts:
first, the check decodes an instruction and determines whether the
instruction ispart of an input, propagation, or assert policy. Second,
the check either updates the map of tainted memory locations, or it
checks to ensure a critical instruction is not acting on tainted data.
Although updating the map of tainted locations and checking
for intrusions is sequential, decoding instructions and determining
whether they apply to our policies can be done in parallel. There-
fore, we divided the taint analysis check into parallel and sequen-
tial phases. Figure 2 shows an example of our parallel taint tracker.
In this example, each numbered box is a single instruction, and
an epoch is two instructions in length (in practice, an epoch en-
compasses thousands of instructions). Within each instrumented
clone, theinstruction is replayed, and then the first phase, called log
preparation, is executed. Log preparation transforms the instruc-
tioninto adependency operation (described in detail in Section6.1)
and places the operation into alog. Thelog of each epoch isdivided
into log segments, which are fixed sized units of memory managed
by Speck. Once the epoch completes, the log is shipped to another
core, where log processing takes place. Log processing processes
each log in sequential order, updates a map of tainted addresses and
registers, and checks for violations of assert policies.
One might imagine that the amount of data generated during log
preparation could be quite large. In fact, our early results showed
that the amount of data was so large that log processing took
longer to complete than log preparation. This problem spurred
the development of dynamic taint compression, which is often
able to reduce the number of operations in a log by a factor of
6. Dynamic taint compression is designed based on two insights:
first, later operations often make earlier operations unnecessary.
Second, the state of the taint map is only important before an assert
check and at the end of each log. We use these observations to
reduce the size of each log in parallel during log preparation. A
detailed explanation appears in Section 6.2. Thus, log preparation
isatwostepprocess. First,loggeneration (describedinSection6.1)
decodes the instructions, maps them to input, propagation, or assert
policies and inserts the operations into the log. Second, dynamic
taint compression (described in Section 6.2) reduces the size of the
log, which accelerates the sequential log processing phase.
Parallel taint analysis implementation
Log generation is implemented as a Pintool that instruments the
targetapplication. Foreachinstruction thatcould potentiallytrigger
an input, propagation, or assert policy, the Pintool writes an 8 byte
log record that describes the operation type and size, the source
location, and the destination location of the instruction.
Each entry in a log may be one of six types. A TAINT record is
insertedwhenaninput policydictatesthat anaddressintheapplica-
tion’s address space is tainted by itsexecution (e.g., TAINT records
are inserted when data is read from a network socket). A CLEAR
record is inserted if an instruction removes taint from a location.
For example, the instruction XOR eax, eax generates a log entry
that clears the eax register of taint. An ASSERT record enforces as-
sert policies by checking to see if a location is tainted. For instance,
ASSERT eax indicates that Speck should halt program execution if
the eax register is tainted.
The remaining record types implement propagation policies.
REPLACE records represent instructions such as MOV that overwrite
the destination operand with the contents of the source operand.
ADD records describe arithmetic and similar operations, where the
destination operand is tainted if either the source or destination is
initiallytainted. SWAP is used for the XCHG instruction, which swaps
thesourceand destinationoperands. Morecomplex operations such
as XADD are described using multiple log records (e.g., a SWAP
followed by an ADD). Operations such as MOVS that move data
between two addresses are encoded as two separate records. Some
instructions do not generate any log records since they do not affect
any taint policy.
During an epoch, multiple log segments are generated to form
a log. Rather than detect the end of a segment by executing an ex-
pensive conditional after inserting each log record, Speck inserts
a guard page at the end of each log segment that is mapped inac-
cessible. After the log segment is filled, writing the next log entry
causes the OS to deliver a signal to the instrumented application.
The signal handler optimizes the filled log segment, as described in
the next section, swaps a fresh log segment into the process address
space, andchanges thepointer of thenext logentry topoint thestart
of the new segment. Finally, Speck reserves a specified amount of
physical memory to store log segments; each log segment is cre-
ated as fixed-sized shared memory segment that is shipped from an
instrumented clones to the log processing program.
Without optimization, it takes longer to process a log than it takes
to run the sequential taint tracker described in Section 5. This is
unfortunate because log processing is fundamentally sequential.
The log processing program must read each log record, decode
the operation, and perform the specified operation. In contrast, the
sequential taint tracker has less overhead — it simply moves or
verifies taint as appropriate for the instruction it is about to execute.
Further, since the sequential taint tracker executes in the address
space of the application it is instrumenting, it enjoys locality of
reference with the addresses it is checking.
Fortunately, we have found that there is substantial room to op-
timize logs after they are generated but before they are processed.
Since Speck speculatively executes instrumented code, it can po-
tentiallyeliminate all taint operations that become moot due to later
execution of theprogram. Incontrast tostaticanalysis toolsthat can
only eliminate a taint operation if it becomes moot along all possi-
ble code paths, the Speck optimizer is dynamic and can eliminate
an operation if it becomes moot along the code path the applica-
tion actually took. In effect, speculative execution allows the paral-
lel taint checker to peer into the future and eliminate work that will
The goal of optimization is two-fold. First, optimization should
be a parallel operation. Second, optimization should derive (and
eliminate) the set of records that, after examining all records in the
log segment, do not have an effect either on the final taint value of
any address or register, oron thetaint value of an address or register
when it is checked for an assert policy violation. For example, if a
register is tainted and later cleared, and no assert policy check on
that register occurred between the two records, the earlier record
can be safely discarded.
Dynamic taint compression is inspired by mark-and-sweep gar-
bage collection. After each log segment is generated, it is then op-
timized independently of other log segments. The optimizer makes
two passes over the log segment. We use a map, similar to the one
described on Section 5, to track the taint values of addresses and
registers. During optimization, each address or register can have
one of three values; either the location is known to be tainted,
known to be free of taint, or its state is unknown. If the state is
unknown, it depends upon the results of prior log segments, and it
might depend upon one ormore recordswithinthecurrent segment.
The second pass examines the map and identifies locations whose
taint value depends upon a list of one or more records within the
Dynamic taint compression
log segment. The optimizer then creates a new log segment, which
is the union of these lists. The second pass also creates a list of
ranges of tainted and free locations. After the second pass, the new,
smaller log segment (and the list of tainted and free locations) is
passed on to the log processing phase.
One can easily imagine more aggressive optimizations. How-
ever, we have found that the dynamic taint compression algorithm
described in this section hits a sweet spot in the design space — it
is very effective (achieving a 6:1 reduction in log size in our exper-
iments), while minimally impacting performance due to its use of
only two sequential passes through the log segment. Investigating
other optimizations is an interesting direction for future work.
The algorithm used in the log processing phase is very similar to
that of the sequential taint tracker. It isimplemented asa standalone
process that reads operations from log segments and uses them to
propagate and check taint, rather than as a Pintool. However, it uses
identical data structures tostore taint data and implements the same
input, propagation, and assert policies.
To process an optimized log segment, the log processing pro-
gram first reads in the set of locations known to be tainted and
the set known to be taint-free. It clears and sets the taint bytes for
these locations accordingly. Speck first orders log segments in the
sequential order they were generated by each instrumented clone.
Second, each log (the set of segments generated by an instrumented
clone) is ordered sequentially as well. Segments from a later epoch
cannot be processed until all segments from all earlier epochs have
been processed. If any ASSERT record operates on tainted data, the
uninstrumented program is halted and an attack reported. Other-
wise, the log processing program blocks until the next sequential
log segment becomes available.
There is an inherent tradeoff in the choice of log segment size.
Larger segment sizes create contention on the memory bus and
pollute the caches on the system. Smaller segment sizes lead to
greater overhead, since a signal handler is executed after each log
is filled. Further, since the log segment is the unit of optimization,
smaller log sizes reduce the efficiency of optimizations that can
be performed. Our parallel taint analysis check uses 512KB logs,
which are small enough to fit in the L2 cache of the computers we
use in our evaluation.
We used two computers in our evaluation. The first computer is an
8-core (dual quad-core) Intel Xeon (5300 series, Core 2 architec-
ture) 2.66GHz workstation with 4GB of memory and a 1.33GHz
bus. Each quad-core chip is made up of 2 dual-core chips stacked
on top of one another. Each dual-core on the chip shares a 4 MB L2
cache. The 8-core computer runs RedHat Enterprise Linux 4 (64
bit) 2.6 kernel. The second computer is a 4-core (dual dual-core)
Intel Xeon (7000 series, NetBurst architecture) 2.8GHz worksta-
tion with 3GB of memory and a 800MHz bus. Each dual-core chip
shares a 2 MB L2 cache. The 4-core computer runs RedHat Enter-
prise 3 (32 bit) 2.4 kernel.
When we evaluated the parallel version of a security check, we
varied the number of cores available to the operating system using
the GRUB boot loader. On the 8-core computer, we could not boot
any 32 bit Linux 2.4 kernel. Since Speculator currently works only
with a 32 bit Linux 2.4 kernel, we could not run speculative execu-
tion for experiments on this machine. The overhead of Speculator
has previously been shown to be quite small, typically adding per-
formance overhead of a few percent orless (Nightingaleet al. 2005,
Number of cores (NetBurst Xeon)
Frames per second
Number of cores (Core 2 Xeon)
Frames per second
This figure shows the frames per second achieved while mplayer decodes an H.264 video clip. Speck parallelizes a check that continuously
scans mplayer’s memory contents. The dashed line shows the performance of the sequential version of the security check. Without any security
check, mplayer decodes 102 fps on the 4-core machine and 175 fps on the 8-core machine. The results represent the mean of 5 trials. Standard
deviations were less than 1%. Note that the scales of the graphs differ.
Figure 3. Process memory analysis: mplayer
Number of cores (NetBurst Xeon)
Transactions per Second
Number of cores (Core 2 Xeon)
Transactions per Second
This figure shows the transactions per-second (TPS) executed by the Postmark benchmark as its files are scanned by a on-access virus scanner
parallelized using Speck. The dashed line shows the performance of the sequential virus scanner. Without any security checks, PostMark
executes 23,992 TPS on the 4-core machine and 53,078 TPS on the 8-core machine. The results represent the mean of 5 trials. Standard
deviations on the 4-core machine were less than 1%. Standard deviations on the 8-core machine for 1-4 cores were less than 1%, deviations for
the remaining cores were: 2.6% for 5 cores, 5.7% for 6-cores 1.8% for 7 cores and 2.8% for 8 cores. Note that the scales of the graphs differ.
Figure 4. On-access virus scanning: PostMark
2006). We confirmed this for Speck by running all benchmarks on
the 4-core machine with and without Speculator.
We first evaluated the performance of the process memory analysis
check, described in Section 4.2, by measuring the rate at which
mplayer can decode an H.264 video clip (a Harry Potter trailer)
while its memory is continuously monitored for a 16-byte item of
sensitive data. Without any security check, mplayer decodes the
movie at a rate of 102 frames per second on the 4-core machine
and 175 frames per second on the 8-core machine.
Our results are shown in Figure 3. With 8 available cores,
mplayer achieves a 7.5x improvement in performance when com-
pared to the sequential version of the check. By using Speck, the
movie is decoded in real-time, which was not possible using the
sequential version of the check. Speck is able to achieve near-ideal
speedups for two reasons. First, the parallel version of the check is
identical to the sequential version, so they perform similarly on one
core. Second, the checks are independent of each other, so perfor-
mance improves linearly with the number of cores.
Process memory analysis
We next measured the performance of the on-access virus scanner,
described inSection4.3, byexecuting thePostMark(Katcher1997)
benchmark. PostMark is an I/O bound benchmark; it was designed
to replicate the small file workloads seen in electronic mail, net-
news, and web-based commerce. We used PostMark version 1.5,
built our virus scanner using ClamAV version 0.91.2 and used a
virus database containing 148,000 signatures. We configured the
virus scanner to scan files on close, emulating a policy that checks
for viruses as the mail or news server writes data to disk. We report
the transactions per-second (TPS) achieved during the benchmark.
With no security checks, PostMark executes 23,992 TPS on the 4-
core machine and 53,078 TPS on the 8-core machine.
Figure 4 shows the TPS executed by PostMark with a single, se-
quential virus scanner, and with multiple virus scanners executing
in parallel with Speck. Since the virus scanners have no dependen-
cies between them, we expected the performance of PostMark to
scale similarly to the memory analysis check. We saw this behavior
on the 4-core machine; compared to the sequential virus scanner,
PostMark is able to execute 3.3x more TPS (on 4 cores) by using
Speck. The 8-core results were quite different. Although initially
Number of cores (NetBurst Xeon)
Frames per second
Number of cores (Core 2 Xeon)
Frames per second
Speck no opt
This figure shows the frames per second achieved while mplayer decodes an H.264 video clip using the taint tracker parallelized with Speck.
The dashed line shows the frames per second using the sequential taint tracker. The solid horizontal line shows the frames per second when
dynamic taint compression is turned off. Without any security check, mplayer decodes 102 frames per second on the 4-core machine and 175
frames per second on the 8-core machine. The results represent the mean of 5 trials. Standard deviations were less than 1%. Note that the
scales of the graphs differ.
Figure 5. Taint analysis: mplayer
scaling well (1.94x more TPS with 2 cores), performance drops
with 4-cores, improves with 5 cores, and then drops again with 7
and 8 available cores. We hypothesize that the single shared mem-
ory bus may be to blame; future architectures with less contention
to main memory may yield better results.
Figure 5 shows the throughput of the parallel and sequential taint
trackers for the mplayer video decoding benchmark. The right
hand graph shows results for the 8-core machine. For one core,
the total computation performed by the parallel version of the
taint tracker is substantially larger than that done by the sequential
version, due to the overhead of generating and optimizing log
segments. However, as the number of available cores increases,
the performance of the parallel version improves. With 4 cores,
the parallel taint tracker outperforms the sequential version. With
8 cores, the parallel version achieves a 2x speedup compared to its
The speedup achievable through parallelization is limitedby the
inherently sequential nature of processing taint propagation. The
8-core graph of Figures 5 shows that optimizing the log before
processing it significantly reduces this bottleneck. Without first
optimizing the log, the time to run the benchmark was limited by
the log processing phase, which runs on a single core. In contrast,
the optimized version of taint analysis shifts enough work to the
parallel phase so that Speck can effectively use additional cores.
Researchers in computer architecture have proposed a style of par-
allelization similar to ours called thread-level speculation (Knight
1986; Sohi et al. 1995; Steffan and Mowry 1998). Thread-level
speculation is intended to take advantage of SMT (simultaneous
multithreading) architectures (Tullsen et al. 1995) by speculatively
scheduling instruction sequences of a single thread in parallel, even
when those sequences may contain data dependencies. Thread-
level speculation requires hardware support to quickly fork a new
thread, to detect data dependencies dynamically, and to discard
the results of incorrect speculative sequences. Two projects have
used the proposed hardware support for thread-level speculation to
parallelize checks that improved the security or reliability of soft-
ware (Oplinger and Lam 2002; Zhou et al. 2004). Oplinger and
Lam (2002) used this style of parallelization to speed up basic-
block profiling, memory access checks, and detection of anomalous
memory loads. Zhou et al. (2004) used this style of parallelization
to speed up functions that monitored memory addresses (watch-
In contrast to prior work on thread-level speculation, our work
shows how to parallelizesecurity checks on commodity processors.
The major challenge without hardware support is amortizing the
high cost of starting new threads (instrumented processes). To ad-
dress this challenge, we start instrumented processes at coarse time
intervals, log non-determinism to synchronize the instrumented
processes with the uninstrumented process, and use OS-level spec-
ulation to enable the uninstrumented process to safely execute be-
yond system calls. OS-level speculation enables speculations to
span a broader set of events and data than thread-level specula-
tion, including interactions between processes and between a pro-
cess and the operating system.
Patil and Fischer (1995) proposed “shadow processing”, which
runs an instrumented process in parallel with an uninstrumented
process. Speck extends this idea to achieve more parallelism by
using speculation to run the uninstrumented process ahead specu-
latively and allowing it to fork multiple instances. A different ap-
proach is to predict which security checks will be needed and to
execute the checks speculatively (Oyama et al. 2005).
OS-level speculation has been used for many purposes besides
security(ChangandGibson1999; FraserandChang 2003; Nightin-
gale et al. 2005; Qin et al. 2005; Li et al. 2005; Nightingale et al.
2006). In the area of security, Anagnostakis et al. (2005) use spec-
ulation to improve security. They test suspicious network input in
an instrumented process (a “shadow honeypot”) and roll it back if
it detects an attack. Speck differs from this work by parallelizing
the work done by security checks, both with the uninstrumented
process and with later security checks.
Researchers at Intel proposed hardware support to enable mon-
itor checks to run in parallel with the original program (Chen et al.
2006). They proposed to modify the processor to trace an unin-
strumented process and ship the trace to a monitor process running
on another core. As with Speck, this enables the uninstrumented
process to run in parallel with the security check. Speck differs
from this work in several ways. Most importantly, Speck enables
a second type of parallelism, which is to run in parallel a sequence
of security checks from different points in a process’s execution.
Second, Speck allows the uninstrumented version to safely execute
system calls that do not externalize output (rather than blocking
on every system call). Third, Speck only ships non-deterministic
events to the monitoring process (rather than a trace of every in-
struction), and this lowers the overhead of synchronization by an
order of magnitude. Finally, Speck requires no hardware support.
Finally, many researchers have explored ways to accelerate spe-
cific security checks. For example, Ho et al. (2006) accelerate
taint analysis by enabling and disabling analysis as taint enters
and leaves a system, and Qin et al. (2006) accelerate taint anal-
ysis through runtime optimizations. These ideas are orthogonal to
our work, which seeks to accelerate arbitrary security checks by
running them in parallel with the original program and with each
Two trends motivate our work on Speck. First, the difficulty of de-
fending systems against attackers isspurring on the development of
active defenses that use expensive run-time security checks. Sec-
ond, future processor improvements will most likely come from
increased parallelism due to multiple cores, rather than from sub-
stantial improvements in clock speed.
Speckstandsat theconfluence of thesetwotrends. Itreduces the
cost of instrumenting security checks by decoupling them from the
critical path of program execution, parallelizing them, and running
them on multiple cores. Our results for three security checks show
substantial speedups on commodity multiprocessors available to-
day. We hope that parallelizing security checks in this manner will
enable the development and use of a new class of powerful security
We thank Mona Attariyan, Ya-Yunn Su, Kaushik Veeraraghavan
and the anonymous reviewers for feedback on this paper. The work
was supported by the National Science Foundation under award
CNS-0509093. Jason Flinnwassupported by NSFCAREERaward
CNS-0346686, Peter Chen was supported by NSF grant CNS-
0614985 and Edmund Nightingale was supported by a Microsoft
ResearchStudent Fellowship. Theviewsand conclusions contained
in this document are those of the authors and should not be inter-
preted as representing the official policies, either expressed or im-
plied, of NSF, Microsoft, the University of Michigan, or the U.S.
ClamAV: The clam anti-virus toolkit, 2007. http://www.clamav.net.
K. G. Anagnostakis, S. Sidiroglou, P. Akritidis, K. Xinidis, E. Markatos,
and A. D. Keromytis. Detecting targeted attacks using shadow honey-
pots. In Proceedings of the 2005 USENIX Security Symposium, August
Thomas C. Bressoud and Fred B. Schneider.
tolerance. In Proceedings of the 15th ACM Symposium on Operating
Systems Principles, pages 1–11, Copper Mountain, CO, December 1995.
Miguel Castro, Manuel Costa, and Tim Harris.
enforcing data-flow integrity. In Proceedings of the 7th Symposium on
Operating Systems Design and Implementation, pages 147–160, Seattle,
WA, October 2006.
Fay Chang and Garth Gibson. Automatic I/Ohint generation through specu-
lative execution. InProceedings of the 3rdSymposium on Operating Sys-
tems Design and Implementation, pages 1–14, New Orleans, LA, Febru-
Shimin Chen, Babak Falsafi, Phillip B. Gibbons, Michael Kozuch, Todd C.
Mowry, Radu Teodorescu, Anastassia Ailamaki, Limor Fix, Gregory R.
Ganger, Bin Lin, and Steven W. Schlosser. Log-Based Architectures for
General-Purpose Monitoring of Deployed Code. 2006 Workshop on Ar-
Securing software by
chitectural and System Support for Improving Software Dependability,
Jim Chow, Ben Pfaff, Tal Garfinkel, and Mendel Rosenblum. Shredding
Your Garbage: Reducing Data Lifetime Through Secure Deallocation.
In Proceedings of the 2005 USENIX Security Symposium, pages 331–
346, August 2005.
Manuel Costa, Jon Crowcroft, Miguel Castro, Antony Rowstron, Lidong
Zhou, Lintao Zhang, and Paul Barham. Vigilante: end-to-end contain-
ment of internet worms. In Proceedings of the 20th ACM Symposium on
Operating Systems Principles, pages 133–147, Brighton, United King-
dom, October 2005.
George W. Dunlap, Samuel T. King, Sukru Cinar, Murtaza A. Basrai, and
Peter M. Chen. ReVirt: Enabling intrusion analysis through virtual-
machine logging and replay. In Proceedings of the 5th Symposium on
Operating Systems Design and Implementation, pages 211–224, Boston,
MA, December 2002.
Keir Fraser and Fay Chang. Operating system I/O speculation: How two
invocations are faster than one. In Proceedings of the 2003 USENIX
Technical Conference, pages 325–338, San Antonio, TX, June 2003.
Alex Ho, Michael Fetterman, Christopher Clark, Andrew Warfield, and
Steven Hand. Practical Taint-based Protection using Demand Emulation.
In Proceedings of EuroSys 2006, 2006.
Jeffrey Katcher. PostMark: A new file system benchmark. Technical Report
TR3022, Network Appliance, 1997.
Tom Knight. An architecture for mostly functional languages. In Proceed-
ings of the 1986 ACM conference on LISP and functional programming,
pages 105–112, 1986.
Calvin Ko and Timothy Redmond. Noninterference and intrusion detection.
In Proceedings of the 2002 IEEE Symposium on Security and Privacy,
pages 177–187, May 2002.
Tong Li, Carla S. Ellis, Alvin R. Lebeck, and Daniel J. Sorin. Pulse: A
dynamic deadlock detection mechanism using speculative execution. In
Proceedings of the 2005 USENIX Technical Conference, 31–44, April
Chi-Keung Luk, Robert Cohn, Robert Muth, Harish Patil, Artur Klauser,
Geoff Lowney, Steven Wallace, Vijay Janapa Reddi, and Kim Hazel-
wood. Pin: Building customized program analysis tools with dynamic
instrumentation. In Programming Language Design and Implementa-
tion, pages 190–200, Chicago, IL, June 2005.
Yevgeniy Miretskiy, Abhijith Das, Charles P. Wright, , and Erez Zadok.
Avfs: An on-access anti-virus file system. In Proceedings of the 13th
USENIX Security Symposium, pages 73–88, 2004.
Nicholas Nethercote and Julian Seward. Valgrind: A framework for heavy-
weight dynamic binary instrumentation. In Proceedings of the Confer-
ence on Programming Language Design and Implementation 2007, San
Diego, CA, June 2007.
JamesNewsome and Dawn Song. Dynamic taint analysis: Automatic detec-
tion, analysis, and signature generation of exploit attacks on commodity
software. In In Proceedings of the 12th Network and Distributed Systems
Security Symposium, February 2005.
Edmund B. Nightingale, Peter M. Chen, and Jason Flinn.
execution in a distributed file system. In Proceedings of the 20th ACM
Symposium on Operating Systems Principles, pages 191–205, Brighton,
United Kingdom, October 2005.
Edmund B. Nightingale, Kaushik Veeraraghavan, Peter M. Chen, and Jason
Flinn.Rethink the sync.In Proceedings of the 7th Symposium on
Operating Systems Design and Implementation, pages 1–14, Seattle,
WA, October 2006.
Jeffrey Oplinger and Monica S. Lam. Enhancing software reliability using
speculative threads. In Proceedings of the 2002 International Confer-
ence on Architectural Support for Programming Languages and Oper-
ating Systems (ASPLOS), pages 184–196, October 2002.
Yoshihiro Oyama, Koichi Onoue, and Akinori Yonezawa.
security checks in sandboxing systems.
International Parallel and Distributed Processing Symposium (IPDPS),
In Proceedings of the 2005
Harish Patil and Charles N. Fischer. Efficient Run-time Monitoring Using Download full-text
Shadow Processing. In Proceedings of the International Workshop on
Automated and Algorithmic Debugging (AADEBUG), pages 119–132,
Niels Provos. Improving host security with system call policies. In Pro-
ceedings of the 12th USENIX Security Symposium, Washington, D.C.,
Feng Qin, Joseph Tucek, Jagadeesan Sundaresan, and Yuanyuan Zhou. Rx:
Treating bugs as allergies—a safe method to survive software failures.
In Proceedings of the 20th ACM Symposium on Operating Systems
Principles, pages 235–248, Brighton, United Kingdom, October 2005.
Feng Qin, Cheng Wang, Zhenmin Li, Ho seop Kim, Yuanyuan Zhou, and
Youfeng Wu. Lift: A low-overhead practical information flow track-
ing system for detecting general security attacks.
nual IEEE/ACM International Symposium on Microarchitecture (MI-
CRO ’06), Orlando, FL, 2006.
Gurindar S. Sohi, Scott E. Breach, and T. N. Vijaykumar.
processors. In Proceedings of the 1995 International Symposium on
Computer Architecture, June 1995.
J. Gregory Steffan and Todd C. Mowry. The potential for using thread-level
data speculation to facilitate automatic parallelization. In Proceedings
of the 1998 Symposium on High Performance Computer Architecture,
In The 39th An-
G. T. Sullivan, D. L. Bruening, I. Baron, T. Garnett, and S. Amarasinghe.
Dynamic native optimization of interpreters.
Workshop on Interpreters, Virtual Machines and Emulators, 2003.
Peter Szor. The Art of Computer Virus Research and Defense. Addison
Dean M. Tullsen, Susan J. Eggers, and Henry M. Levy.
multithreading: Maximizing on-chip parallelism. In Proceedings of the
1995 International Symposium on Computer Architecture, June 1995.
David Wagner and Drew Dean.Intrusion detection via static analysis.
In Proceedings of 2001 IEEE Symposium on Computer Security and
Privacy2001, pages 156–169, 2001.
Min Xu, Rastislav Bodik, and Mark D. Hill. A Flight Data Recorder for
Enabling Full-system Multiprocessor Deterministic Replay. In Proceed-
ings of the 2003 International Symposium on Computer Architecture,
Pin Zhou, Feng Qin, Wei Liu, Yuanyuan Zhou, and Josep Torrellas.
iWatcher: Efficient architectural support for software debugging. In Pro-
ceedings of the 2004 International Symposium on Computer Architec-
ture, June 2004.
In Proceedings of the