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Robust and Efficient Elimination of Cache and
Timing Side Channels

Benjamin A. Braun1, Suman Jana1, and Dan Boneh1

1Stanford University

Abstract—Timing and cache side channels provide powerful
attacks against many sensitive operations including cryptographic
implementations. Existing defenses cannot protect against all
classes of such attacks without incurring prohibitive performance
overhead. A popular strategy for defending against all classes
of these attacks is to modify the implementation so that the
timing and cache access patterns of every hardware instruction
is independent of the secret inputs. However, this solution is
architecture-specific, brittle, and difficult to get right. In this
paper, we propose and evaluate a robust low-overhead technique
for mitigating timing and cache channels. Our solution requires
only minimal source code changes and works across multiple lan-
guages/platforms. We report the experimental results of applying
our solution to protect several C, C++, and Java programs. Our
results demonstrate that our solution successfully eliminates the
timing and cache side-channel leaks while incurring significantly
lower performance overhead than existing approaches.

I. INTRODUCTION

Defending against cache and timing side channel attacks
is known to be a hard and important problem. Timing and
cache attacks can be used to extract cryptographic secrets
from running systems [14, 15, 23, 29, 35, 36, 36, 40], spy
on Web user activity [12], and even undo the privacy of
differential privacy systems [5, 24]. Attacks exploiting timing
side channels have been demonstrated for both remote and
local adversaries. A remote attacker is separated from its target
by a network [14, 15, 29, 36] while a local attacker can execute
unprivileged spyware on the target machine [7, 9, 11, 36, 45,
47].

Most existing defenses against cache and timing attacks
only protect against a subset of attacks and incur significant
performance overheads. For example, one way to defend
against remote timing attacks is to make sure that the timing of
any externally observable events are independent of any data
that should be kept secret. Several different strategies have
been proposed to achieve this, including application-specific
changes [10, 27, 30], static transformation [17, 20], and
dynamic padding [6, 18, 24, 31, 47]. However, none of these
strategies defend against local timing attacks where the attacker
spies on the target application by measuring the target’s impact
on the local cache and other resources. Similarly, the strategies
for defending against local cache attacks like static partitioning
of resources [28, 37, 43, 44], flushing state [50], obfuscating
cache access patterns [9, 10, 13, 35, 40], and moderating
access to fine-grained timers [33, 34, 42], also incur significant
performance penalties while still leaving the target potentially
vulnerable to timing attacks. We survey these methods in
related work (Section VIII).

A popular approach for defending against both local and
remote timing attacks is to ensure that the low-level instruction
sequence does not contain instructions whose performance
depends on secret information. This can be enforced by
manually re-writing the code, as was done in OpenSSL1, or by
changing the compiler to ensure that the generated code has
this property [20].

Unfortunately, this popular strategy can fail to ensure
security for several reasons. First, the timing properties of
instructions may differ in subtle ways from one architecture
to another (or even from one processor model to another)
resulting in an instruction sequence that is unsafe for some
architectures/processor models. Second, this strategy does not
work for languages like Java where the Java Virtual Machine
(JVM) optimizes the bytecode at runtime and may inad-
vertently introduce secret-dependent timing variations. Third,
manually ensuring that a certain code transformation prevents
timing attacks can be extremely difficult and tedious, as was
the case when updating OpenSSL to prevent the Lucky-thirteen
timing attack [32].

Our contribution. We propose the first low-overhead,
application-independent, and cross-language defense that can
protect against both local and remote timing attacks with
minimal application code changes. We show that our defense
is language-independent by applying the strategy to protect
applications written in Java and C/C++. Our defense requires
relatively simple modifications to the underlying OS and can
run on off-the-shelf hardware.

We implement our approach in Linux and show that the
execution times of protected functions are independent of
secret data. We also demonstrate that the performance overhead
of our defense is low. For example, the performance overhead
to protect the entire state machine running inside a SSL/TLS
server against all known timing- and cache-based side channel
attacks is less than 5% in connection latency.

We summarize the key insights behind our solution (de-
scribed in detail in Section IV) below.

• We leverage programmer code annotations to identify
and protect sensitive code that operates on secret data.
Our defense mechanism only protects the sensitive func-
tions. This lets us minimize the performance impact of
our scheme by leaving the performance of non-sensitive
functions unchanged.

1In the case of RSA private key operations, OpenSSL uses an additional
defense called blinding.

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• We further minimize the performance overhead by sepa-
rating and accurately accounting for secret-dependent and
secret-independent timing variations. Secret-independent
timing variations (e.g., the ones caused by interrupts, the
OS scheduler, or non-secret execution flow) do not leak
any sensitive information to the attacker and thus are
treated differently than secret-dependent variations by our
scheme.

• We demonstrate that existing OS services like schedulers
and hardware features like memory hierarchies can be
leveraged to create a lightweight isolation mechanism that
can protect a sensitive function’s execution from other
local untrusted processes and minimize timing variations
during the function’s execution.

• We show that naive implementations of delay loops in
most existing hardware leak timing information due to
the underlying delay primitive’s (e.g., NOP instruction)
limited accuracy. We create and evaluate a new scheme
for implementing delay loops that prevents such leakage
while still using existing coarse-grained delay primitives.

• We design and evaluate a lazy state cleansing mechanism
that clears the sensitive state left in shared hardware
resources (e.g., branch predictors, caches, etc.) before
handing them over to an untrusted process. We find that
lazy state cleansing incurs significantly less overhead than
performing state cleaning as soon as a sensitive function
finishes execution.

II. KNOWN TIMING ATTACKS

Before describing our proposed defense we briefly survey
different types of timing attackers. In the previous section, we
discussed the difference between a local and a remote timing
attacker: a local timing attacker, in addition to monitoring the
total computation time, can spy on the target application by
monitoring the state of shared hardware resources such as the
local cache.

Concurrent vs. non-concurrent attacks. In a concurrent
attack, the attacker can probe shared resources while the target
application is operating. For example, the attacker can measure
timing information or inspect the state of the shared resources
at intermediate steps of a sensitive operation. The attacker’s
process can control the concurrent access by adjusting its
scheduling parameters and its core affinity in the case of
symmetric multiprocessing (SMP).

A non-concurrent attack is one in which the attacker only
gets to observe the timing information or shared hardware state
at the beginning and the end of the sensitive computation.
For example, a non-concurrent attacker can extract secret
information using only the aggregate time it takes the target
application to process a request.

Local attacks. Concurrent local attacks are the most prevalent
class of timing attacks in the research literature. Such attacks
are known to be able to extract the secret/private key against
a wide-range of ciphers including RSA [4, 36], AES [23, 35,
40, 46], and ElGamal [49]. These attacks exploit information
leakage through a wide range of shared hardware resources: L1
or L2 data cache [23, 35, 36, 40], L3 cache [26, 46], instruction
cache [1, 49], branch predictor cache [2, 3], and floating-point
multiplier [4].

There are several known local non-concurrent attacks as
well. Osvik et al. [35], Tromer et al. [40], and Bonneau
and Mironov [11] present two types of local, non-concurrent
attacks against AES implementations. In the first, prime and
probe, the attacker “primes” the cache, triggers an AES en-
cryption, and “probes” the cache to learn information about the
AES private key. The spy process primes the cache by loading
its own memory content into the cache and probes the cache by
measuring the time to reload the memory content after the AES
encryption has completed. This attack involves the attacker’s
spy process measuring its own timing information to indirectly
extract information from the victim application. Alternatively,
in the evict and time strategy, the attacker measures the time
taken to perform the victim operation, evicts certain chosen
cache lines, triggers the victim operation and measure its
execution time again. By comparing these two execution times,
the attacker can find out which cache lines were accessed
during the victim operation. Osvik et al. were able to extract
an 128-bit AES key after only 8,000 encryptions using the
prime and probe attack.

Remote attacks. All existing remote attacks [14, 15, 29, 36]
are non-concurrent, however this is not fundamental. A hy-
pothetical remote, yet concurrent, attack would be one in
which the remote attacker submits requests to the victim
application at the same time that another non-adversarial client
sends some requests containing sensitive information to the
victim application. The attacker may then be able to measure
timing information at intermediate steps of the non-adversarial
client’s communication with the victim application and infer
the sensitive content.

III. THREAT MODEL

We allow the attacker to be local or remote and to execute
concurrently or non-concurrently with the target application.
We assume that the attacker can only run spy processes as
a different non-privileged user (i.e., no super-user privileges)
than the owner of the target application. We also assume
that the spy process cannot bypass the standard user-based
isolation provided by the operating system. We believe that
these are very realistic assumptions because if either one of
these assumptions fail, the spy process can steal the user’s
sensitive information without resorting to side channel attacks
in most existing operating systems.

In our model, the operating system and the underlying
hardware are trusted. Similarly, we expect that the attacker
does not have physical access to the hardware and cannot
monitor side channels such as electromagnetic radiations,
power use, or acoustic emanations. We are only concerned
with timing and cache side channels since they are the easiest
side channels to exploit without physical access to the victim
machine.

IV. OUR SOLUTION

In our solution, developers annotate the functions perform-
ing sensitive computation(s) that they would like to protect.
For the rest of the paper, we refer to such functions as
protected functions. Our solution instruments the protected
functions such that our stub code is invoked before and after
execution of each protected function. The stub code ensures

that the protected functions, all other functions that may be
invoked as part of their execution, and all the secrets that they
operate on are safe from both local and remote timing attacks.
Thus, our solution automatically prevents leakage of sensitive
information by all functions (protected or unprotected) invoked
during a protected function’s execution.

Our solution ensures the following properties for each
protected function:

• We ensure that the execution time of a protected function
as observed by either a remote or a local attacker is
independent of any secret data the function operates on.
This prevents an attacker from learning any sensitive in-
formation by observing the execution time of a protected
function.

• We clean any state left in the shared hardware resources
(e.g., caches) by a protected function before handing
the resources over to an untrusted process. As described
earlier in our threat model (Section III), we treat any
process as untrusted unless it belongs to the same user
who is performing the protected computation. We cleanse
shared state only when necessary in a lazy manner to
minimize the performance overhead.

• We prevent other concurrent untrusted processes from ac-
cessing any intermediate state left in the shared hardware
resources during the protected function’s execution. We
achieve this by efficiently dynamic partitioning the shared
resources while incurring minimal performance overhead.

L2#cache#

L3#cache#

L1#cache#

L2#cache#

L1#cache#

L2#cache#

L1#cache#

per,user#page#coloring#isolates#protected#
func7on’s#cache#lines##

•  no#user#process#can#preempt#protected#func7ons#
•  apply#padding#to#make#7ming#secret,independent#
•  lazily#clean#per,core#resources#

core#1# core#2# core#3#
protected#
func7on#

untrusted#
process#

untrusted#
process#

Fig. 1: Overview of our solution

Figure 1 shows the main components of our solution.
We use two high-level mechanisms to provide the properties
described above for each protected function: time padding and
preventing leakage through shared resources. We first briefly
summarize these mechanisms below and then describe them
in detail in Sections IV-A and IV-B.

Time padding. We use time padding to make sure that
a protected function’s execution time does not depend on

the secret data. The basic idea behind time padding is sim-
ple—pad the protected function’s execution time to its worst-
case runtime over all possible inputs. The idea of padding
execution time to an upper limit to prevent timing channels
itself is not new and has been explored in several prior
projects [6, 18, 24, 31, 47]. However, all these solutions
suffer from two major problems which prevent them from
being adopted in real-world setting: i) they incur prohibitive
performance overhead (90−400% in macro-benchmarks [47])
because they have to add a large amount of time padding in
order to prevent any timing information leakage to a remote
attacker, and ii) they do not protect against local adversaries
who can infer the actual unpadded execution time through side
channels beyond network events (e.g., by monitoring the cache
access patterns at periodic intervals).

We solve both of these problems in this paper. One of our
main contributions is a new low-overhead time padding scheme
that can prevent timing information leakage of a protected
function to both local and remote attackers. We minimize
the required time padding without compromising security by
adapting the worst-case time estimates using the following
three principles:

1) We adapt the worst-case execution estimates to the target
hardware and the protected function. We do so by pro-
viding an offline profiling tool to automatically estimate
worst-case runtime of a particular protected function
running on a particular target hardware platform. Prior
schemes estimate the worst-case execution times for
complete services (i.e., web servers) across all possible
hardware configurations. This results in an over-estimate
of the time pad that hurts performance.

2) We protect against local (and remote) attackers by ensur-
ing that an untrusted process cannot intervene during a
protected function’s execution. We apply time padding at
the end of every protected function’s execution. This en-
sures minimal overhead while preventing a local attacker
from learning the running time of protected functions.
Prior schemes applied a large time pad before sending a
service’s output over the network. Such schemes are not
secure against local attackers who can use local resources,
such as cache behavior, to infer the execution time of
individual protected functions.

3) Timing variations result from many factors. Some are
secret-dependent and must be prevented, while others
are secret independent and cause no harm. For example,
timing variations due to the OS scheduler and interrupt
handlers are generally harmless. We accurately measure
and account for secret-dependent variations and ignore
the secret-independent variations. This lets us compute an
optimal time pad needed to protect secret data. None of
the existing time padding schemes distinguish between
the secret-dependent and secret-independent variations.
This results in unnecessarily large time pads, even when
secret-dependent timing variations are small.

Preventing leaks via shared resources. We prevent in-
formation leakage through shared resources without adding
significant performance overhead to the process executing the
protected function or to other (potentially malicious) processes.
Our approach is as follows:

• We leverage the multi-core processor architecture found
in most modern processors to minimize the amount of
shared resources during a protected function’s execution
without hurting performance. We dynamically reserve
exclusive access to a physical core (including all per-
core caches such as L1 and L2) while it is executing
a protected function. This ensures that a local attacker
does not have concurrent access to any per-core resources
while a protected function is accessing them.

• For L3 caches shared across multiple cores, we use page
coloring to ensure that cache accesses during a protected
function’s execution are restricted within a reserved por-
tion of the L3 cache. We further ensure that this reserved
portion is not shared with other users’ processes. This
prevents the attacker from learning any information about
protected functions through the L3 cache.

• We lazily cleanse the state left in both per-core resources
(e.g., L1/L2 caches, branch predictors) and resources
shared across cores (e.g., L3 cache) only before handing
them over to untrusted processes. This minimizes the
overhead caused by the state cleansing operation.

A. Time padding

We design a safe time padding scheme that defends against
both local and remote attackers inferring sensitive information
from observed timing behavior of a protected function. Our de-
sign consists of two main components: estimating the padding
threshold and applying the padding safely without leaking any
information. We describe these components in detail next.

Determining the padding value. Our time padding only
accounts for secret-dependent time variations. We discard
variations due to interrupts or OS scheduler preemptions. To
do so we rely Linux’s ability to keep track of the number of
external preemptions. We adapt the total padding time based
on the amount of time that a protected function is preempted
by the OS.

• Let Tmax be the worst-case execution time of a protected
function when no external preemptions occur.

• Let Text preempt be the worst-case time spent during pre-
emptions given the set of n preemptions that occur during
the execution of the protected function.

Our padding mechanism pads the execution of each protected
function to Tpadded cycles, where

Tpadded = Text preempt + Tmax.

This leaks the amount of preemption time to the attacker,
but nothing else. Since this is independent of the secret, the
attacker learns nothing useful.

Estimating Tmax. Our time padding scheme requires a tight
estimate of the worst-case execution time (WCET) of every
protected function. There are several prior projects that try
to estimate WCET through different static analysis tech-
niques [19, 25]. However, these techniques require precise and
accurate models of the target hardware (e.g., cache, branch
target buffers, etc.) which are often very hard to get in practice.
In our implementation we use a simple dynamic profiling
method to estimate WCET described below. Our time padding

time

Padding target:

Leak

Fig. 2: Time leakage due to naive padding

scheme is not tied to any particular WCET estimation method
and can work with other estimation tools.

We estimate the WCET, Tmax, through dynamic offline pro-
filing of the protected function. Since this value is hardware-
specific, we perform the profiling on the actual hardware
that will run protected functions. To gather profiling informa-
tion, we run an application that invokes protected functions
with an input generating script provided by the application
developer/system administrator. To reduce the possibility of
overtimes occurring due to uncommon inputs, it is important
that the script generate both common and uncommon inputs.
We instrument the protected functions in the application so
that the worst-case performance behavior is stored in a profile
file. We compute the padding parameters based on the profiling
results.

To be conservative, we obtain all profiling measurements
for the protected functions under high load conditions (i.e., in
parallel with other application that produces significant loads
on both memory and CPU). We compute Tmax from these
measurements such that it is the worst-case timing bound when
at most a κ fraction of all profiling readings are excluded. κ is a
security parameter which provides a tradeoff between security
and performance. Higher values of κ reduce Tmax but increase
the chance of overtimes. For our prototype implementation we
set κ to 10−5.

Safely applying padding. Once the padding amount has been
determined using the techniques described earlier, waiting for
the target amount might seem easy at first glance. However,
there are two major issues that make application of padding
complicated in practice as described below.

Handling limited accuracy of padding loops. As our solution
depends on fine-grained padding, a naive padding scheme may
leak information due to limited accuracy of any padding loops.
Figure 2 shows that a naive padding scheme that repeatedly
measures the elapsed time in a tight loop until the target time
is reached leaks timing information. This is because the loop
can only break when the condition is evaluated, and hence
if one iteration of the loop takes u cycles then the padding
loop leaks timing information mod u. Note that earlier timing
padding schemes do not get affected by this problem as their
padding amounts are significantly larger than ours.

Our solution guarantees that the distribution of running
times of a protected function for some set of private inputs
is indistinguishable from the same distribution produced when
a different set of private inputs to the function are used. We

call this property the safe padding property. We overcome
the limitations of the simple wait loop by performing a
timing randomization step before entering the simple wait
loop. During this step, we perform m rounds of a randomized
waiting operation. This goal of this step is to ensure that the
amount of time spent in the protected function before the
beginning of the simple wait loop, when taken modulo u, the
stable period of the simple timing loop (i.e. disregarding the
first few iterations), is close to uniform. This technique can be
viewed as performing a random walk on the integers modulo u
where the runtime distribution of the waiting operation is the
support of the walk and m is the number of steps walked. Prior
work by Chung et al. [16] has explored the sufficient conditions
for the number of steps in a walk and its support that produce
a distribution that is exponentially close to uniform.

For the purposes of this paper, we perform timing random-
ization using a randomized operation with 256 possible inputs
that runs for X + c cycles on input X where c is a constant.
We describe the details of this operation in Section V. We
then choose m to defeat our empirical statistical tests under
pathological conditions that are very favorable to an attacker
as shown in Section VI.

For our scheme’s guarantees to hold, the randomness used
inside the randomized waiting operation must be generated
using a cryptographically secure generator. Otherwise, if an
attacker can predict the added random noise, she can subtract
it from the observed padded time and hence derive the original
timing signal, modulo u.

A padding scheme that pads to the target time Tpadded
using a simple padding loop and performs the randomization
step after the execution of the protected function will not
leak any information about the duration of the protected
function, as long as the following conditions hold: (i) no
preemptions occur; (ii) the randomization step successfully
yields a distribution of runtimes that is uniform modulo u;
(iii) The simple padding loop executes for enough iterations
so that it reaches its stable period. The security of this scheme
under these assumptions can be proved as follows.

Let us assume that the last iteration of the simple wait
loop take u cycles. Assuming the simple wait loop has iterated
enough times to reach its stable period, we can safely assume
that u does not depend on when the simple wait loop started
running. Now, due to the randomization step, we assume that
the amount of time spent up to the start of the last iteration of
the simple wait loop, taken modulo u, is uniformly distributed.
Hence, the loop will break at a time that is between the
target time and the target time plus u − 1. Because the last
iteration began when the elapsed execution time was uniformly
distributed modulo u, these u different cases will occur with
equal probability. Hence, regardless of what is done within
the protected function, the padded duration of the function
will follow a uniform distribution of u different values after
the target time. Therefore, the attacker will not learn anything
from observing the padded time of the function.

To reduce the worst-case performance cost of the random-
ization step, we generate the required randomness at the start
of the protected function, before measuring the start time of
the protected function. This means that any variability in the
runtime of the randomness generator does not increase Tpadded.

// At the return point of a protected function:
// Tbegin holds the time at function start
// Ibegin holds the preemption count at function start

1 for j = 1 to m
2 Short-Random-Delay()
3 Ttarget = Tbegin + Tmax
4 overtime = 0
5 for i = 1 to ∞
6 before = Current-Time()
7 while Current-Time() < Ttarget, re-check.
8 // Measure preemption count and adjust target
9 Text preempt = (Preemptions()− Ibegin) ·Tpenalty

10 Tnext = Tbegin + Tmax + Text preempt + overtime
11 // Overtime-detection support
12 if before ≥ Tnext and overtime = 0
13 overtime = Tovertime
14 Tnext = Tnext + overtime
15 // If no adjustment was made, break
16 if Tnext = Ttarget
17 return
18 Ttarget = Tnext

Fig. 3: Algorithm for applying time padding to a protected function’s
execution.

Handling preemptions occurring inside the padding loop.
The scheme presented above assumes that no external pre-
emptions can occur during the the execution of the padding
loop itself. However, blocking all preemptions during the
padding loop will degrade the responsiveness of the system. To
avoid such issues, we allow interrupts to be processed during
the execution of the padding loop and update the padding
time accordingly. We repeatedly update the padding time in
response to preemptions until a “safe exit condition” is met
where we can stop padding.

Our approach is to initially pad to the target value Tpadded,
regardless of how many preemptions occur. We then repeatedly
increase Text preempt and pad to the new adjusted padding target
until we execute a padding loop where no preemptions occur.
The pseudocode of our approach is shown in Figure 3. Our
technique does not leak any information about the actual
runtime of the protected function as the final padding target
only depends on the pattern of preemptions but not on the
initial elapsed time before entering the padding loops. Note
that forward progress in our padding loops is guaranteed as
long as preemptions are rate limited on the cores executing
protected functions.

The algorithm computes Text preempt based on observed
preemptions simply by multiplying a constant Tpenalty by the
number of preemptions. Since Text preempt should match the
worst-case execution time of the observed preemptions, Tpenalty
is the worst-case execution time of any single preemption.
Like Tmax, Tpenalty is machine specific and can be determined
empirically from profiling data.

Handling overtimes. Our WCET estimator may miss a
pathological input that causes the protected function to run for
significantly more time than on other inputs. While we never

observed this in our experiments, if such a pathological input
appeared in the wild, the protected function may take longer
than the estimated worst-case bound and this will result in
an overtime. This leaks information: the attacker learns that a
pathological input was just processed. We therefore augment
our technique to detect such overtimes, i.e., when the elapsed
time of the protected function, taking interrupts into account,
is greater than Tpadded.

One option to limit leakage when such overtimes are
detected is to refuse to service such requests. The system
administrator can then act by either updating the secrets (e.g.,
secret keys) or increasing the parameter Tmax of the model.

We also support updating Tmax of a protected function
on the fly without restarting the running application. The
padding parameters are stored in a file that has the same
access permissions as the application/library containing the
protected function. This file is memory-mapped when the
corresponding protected function is called for the first time.
Any changes to the memory-mapped file will immediately
impact the padding parameters of all applications invoking the
protected function unless they are in the middle of applying
the estimated padding.

Note that each overtime can at most leak log(N) bits of
information, where N is the total number of timing measure-
ments observed by the attacker. To see why, consider a string
of N timing observations made by an attacker with at most
B overtimes. There can be < NB such unique strings and
thus the maximum information content of such a string is
< Blog(N) bits, i.e., < log(N) bits per overtime. However,
the actual effect of such leakage depends on how much entropy
an application’s timing patterns for different inputs have. For
example, if an application’s execution time for a particular
secret input is significantly larger than all other inputs, even
leaking 1 bit of information will be enough for the attacker to
infer the complete secret input.

Minimizing external preemptions. Note that even though
Tpadded does not leak any sensitive information, padding to this
value will incur significant performance overhead if Text preempt
is high due to frequent or long-running preemptions during
the protected function’s execution. Therefore, we minimize the
external events that can delay the execution of a protected
function. We describe the main external sources of delays and
how we deal with them in detail below.

• Preemptions by other user processes. Under regular
circumstances, execution of a protected function may
be preempted by other user processes. This can delay
the execution of the protected function as long as the
process is preempted. Therefore, we need to minimize
such preemptions while still keeping the system usable.
In our solution, we prevent preemptions by other user
processes during the execution of a protected function
by using a scheduling policy that prevents migrating
the process to a different core and prevents other user
processes from being scheduled on the same core during
the duration of the protected function’s execution.

• Preemptions by interrupts. Another common source
of preemption is the hardware interrupts served by the
core executing a protected function. One way to solve

this problem is to block or rate limit the number of
interrupts that can be served by a core while executing a
protected function. However, such a technique may make
the system non-responsive under heavy load. For this
reason, in our current prototype solution, we do not apply
such techniques.
Note that some of these interrupts (e.g., network inter-
rupts) can be triggered by the attacker and thus can be
used by the attacker to slow down the protected function’s
execution. However, in our solution, such an attack in-
creases Text preempt, and hence degrades performance, but
does not cause information leakage.

• Paging. An attacker can potentially arbitrarily slow down
the protected function by causing memory paging events
during the execution of a protected function. To avoid
such cases, our solution forces each process executing a
protected function to lock all of its pages in memory and
disables page swapping. As a consequence, our solution
currently does not allow processes that allocate more
memory than is physically available in the target system
to use protected functions.

• Hyperthreading. Hyperthreading is a technique sup-
ported by modern processor cores where one physical
core supports multiple logical cores. The operating system
can independently schedule tasks on these logical cores
and the hardware transparently takes care of sharing the
underlying physical core. We observed that protected
functions executing on a core with hyperthreading enabled
can encounter large amounts of slowdown. This slowdown
is caused because the other concurrent processes execut-
ing on the same physical core can interfere with access
to some of the CPU resources.
One potential way of avoiding this slowdown is to con-
figure the OS scheduler to prevent any untrusted process
from running concurrently on a physical core with a
process in the middle of a protected function. However,
such a mechanism may result in high overheads due
to the cost of actively unscheduling/migrating a process
running on a virtual core. For our current prototype
implementation, we simply disable hyperthreading as part
of system configuration.

• CPU frequency scaling. Modern CPUs include mech-
anisms to change the operating frequency of each core
dynamically at runtime depending on the current work-
load to save power. If a core’s frequency decreases in
the middle of the execution of a protected function or it
enters the halt state to save power, it will take longer in
real-time, increasing Tmax. To reduce such variations, we
disable CPU frequency scaling and low-power CPU states
when a core executes a protected function.

B. Preventing leakage through shared resources

We prevent information leakage from protected functions
through shared resources in two ways: isolating shared re-
sources from other concurrent processes and lazily cleansing
state left in shared resources before handing them over to other
untrusted processes. Isolating shared resources of protected
functions from other concurrent processes help in preventing
local timing and cache attacks as well as improving perfor-
mance by minimizing variations in the runtime of protected

functions.

Isolating per-core resources. As described earlier in Sec-
tion IV-A, we disable hyperthreading on a core during a
protected function’s execution to improve performance. This
also ensures that an attacker cannot run spy code that snoops
on per-core state while a protected function is executing. We
also prevent preemptions from other user processes during the
execution of protected function and thus ensure that the core
and its L1/L2 caches are dedicated for the protected function.

Preventing leakage through performance counters. Modern
hardware often contain performance counters that keep track of
different performance events such as the number of cache evic-
tions or branch mispredictions occurring on a particular core.
A local attacker with access to these performance counters may
infer the secrets used during a protected function’s execution.
Our solution, therefore, restricts access to performance mon-
itoring counters so that a user’s process cannot see detailed
performance metrics of another user’s processes. We do not
restrict, however, a user from using hardware performance
counters to measure the performance of their own processes.

Preventing leakage through L3 cache. As L3 cache is a
shared resources across multiple cores, we use page coloring
to dynamically isolate the protected function’s data in the L3
cache. To support page coloring we modify the OS kernel’s
physical page allocators so that they do not allocate pages
having any of C reserved secure page colors, unless the caller
specifically requests a secure color. Pages are colored based
on which L3 cache sets a page maps to. Therefore, two pages
with different colors are guaranteed never to conflict in the L3
cache in any of their cache lines.

In order to support page coloring, we disable transparent
huge pages and set up access control to huge pages. An
attacker that has access to a huge page can evade the isolation
provided by page coloring, since a huge page can span
multiple page colors. Hence, we prevent access to huge pages
(transparently or by request) for non-privileged users.

As part of our implementation of page coloring, we also
disable memory deduplication features, such as kernel same-
page merging. This prevents a secure-colored page mapped
into one process from being transparently mapped as shared
into another process. Disabling memory deduplication is not
unique to our solution and has been used in the past in
hypervisors to prevent leakage of information across different
virtual machines [39].

When a process calls a protected function for the first time,
we invoke a kernel module routine to remap all pages allocated
by the process in private mappings (i.e., the heap, stack, text-
segment, library code, and library data pages) to pages that
are not shared with any other user’s processes. We also ensure
these pages have a page color reserved by the user executing
the protected function. The remapping transparently changes
the physical pages that a process accesses without modifying
the virtual memory addresses, and hence requires no special
application support. If the user has not yet reserved any page
colors or there are no more remaining pages of any of her
reserved page colors, the OS allocates one of the reserved
colors for the user. Also, the process is flagged with a ”secure-
color” bit. We modify the OS so that it recognizes this flag and

ensures that the future pages allocated to a private mapping for
the process will come from a reserved page color for the user.
Note that since we only remap private mappings, we do not
protect applications that access a shared mapping from inside
a protected function.

This strategy for allocating page colors to users has a minor
potential downside that such a system restricts the numbers of
different users’ processes that can concurrently call protected
functions. We believe that such a restriction is a reasonable
trade-off between security and performance.

Lazy state cleansing. To ensure that an attacker does not
see the tainted state in a per-core resource after a protected
function finishes execution, we lazily delete all per core re-
sources. When a protected function returns, we mark the CPU
as “tainted” with the user ID of the caller process. The next
time the OS attempts to schedule a process from a different
user on the core, it will first flush all per-CPU caches, including
the L1 instruction cache, L1 data cache, L2 cache, Branch
Translation Buffer (BTB), and Translation lookaside buffer
(TLB). Such a scheme ensures that the overhead of flushing
these caches can be amortized over multiple invocations of
protected functions by the same user.

V. IMPLEMENTATION

We built a prototype implementation of our protection
mechanism for a system running Linux OS. We describe the
different components of our implementation below.

A. Programming API

We implement a new function annotation FIXED TIME
for the C/C++ language that indicates that a function should
be protected. The annotation can be specified either in the
declaration of the function or at its definition. Adding this
annotation is the only change to a C/C++ code base that a
programmer has to make in order to use our solution. We
wrote a plugin for the Clang C/C++ compiler that handles
this annotation. The plugin automatically inserts a call to
the function fixed time begin at the start of the protected
function and a call to fixed time end at any return point of
the function. These functions protect the annotated function
using the mechanisms described in Section IV.

Alternatively, a programmer can also call these functions
explicitly. This is useful for protecting ranges of code within
function such as the state transitions of the TLS state machine
(see Section VI-B). We provide a Java native interface wrapper
to both fixed time begin and fixed time end functions, for
supporting protected functions written in Java.

B. Time padding

For implementing time padding loops, we read from the
timestamp counter in x86 processors to collect time measure-
ments. In most modern x86 processors, including the one we
tested on, the timestamp counter has a constant frequency
regardless of the power saving state of a processor. We generate
pseudorandom bytes for the randomized padding step using
the ChaCha/8 stream cipher [8]. We use a value of 300 µs
for Tpenalty as this bounds the worst-case slowdown due to a
single interrupt we observed in our experiments.

Our implementation of the randomized wait operation takes
an input X and simply performs X +c noops in a loop, where
c is a large enough value so that the loop takes one cycle
longer for each additional iteration. We observe that c = 46 is
sufficient to achieve this property.

Some of the OS modifications specified in our solution
are implemented as a loadable kernel module. This module
supports an IOCTL call to mark a core as tainted at the
end of a protected function’s execution. The module also
supports an IOCTL call that enables fast access to the interrupt
and context-switch count. In the standard Linux kernel, the
interrupt count is usually accessed through the proc file system
interface. However, such an interface is too slow for our
purposes. Instead, our kernel module allocates a page of
counters that is mapped into the virtual address space of the
calling process. The task struct of the calling process also
contains a pointer to these counters. We modify the kernel
to check on every interrupt and context switch if the current
task has such a page, and if so, to increment the corresponding
counter in that page.

Offline profiling. We provide a profiling wrapper script,
fixed time record . sh, that computes worst-case execution
time parameters of each protected function as well as the
worst-case slowdown on that function due to preemptions by
different interrupts or kernel tasks.

The profiling script automatically generates profiling in-
formation for all protected functions in an executable by
running the application on different inputs. During the pro-
filing process, we run a variety of applications in parallel
to create a stress-testing environment that triggers worst-case
performance of the protected function. To allow the stress
testers to maximally slow down the user application, we reset
the scheduling parameters and CPU affinity of a thread at
the start and end of every protected function. One stress
tester generates interrupts at a high frequency using a simple
program that generates a flood of UDP packets to the loopback
network interface. We also run the mprime2, systester3, and
the LINPACK benchmark4 to cause high CPU load and large
amounts of memory contention.

C. Prevent leakage through shared resources

Isolating a processor core and core-specific caches. We
disable hyperthreading in Linux by selectively disabling virtual
cores. This prevents any other processes from interfering with
the execution of a protected function. As part of our prototype,
we also implement a simple version of the page coloring
scheme described in Section IV.

We prevent a user from observing hardware performance
counters showing the performance behavior of other users’
processes. The perf events framework on Linux mediates
access to hardware performance counters. We configure the
framework to allow accessing per-CPU performance counters
only by the privileged users. Note that an unprivileged user can

2http://www.mersenne.org/
3http://systester.sourceforge.net
4https://software.intel.com/en-us/articles/intel-math-kernel-library-linpack-

download/

still access per-process performance counters that measure the
performance of their own processes.

For ensuring that a processor core executing a pro-
tected function is not preempted by other user processes,
as specified in Section IV, we depend on a scheduling
mode that prevents other userspace processes from preempting
a protected function. For this purpose, we use the Linux
SCHED FIFO scheduling mode at maximum priority. In order
to be able to do this, we allow unprivileged users to use
SCHED FIFO at priority 99 by changing the limits in the
/etc/security/limits.conf file.

One side effect of this technique is that if a protected
function manually yields to the scheduler or perform blocking
operations, the process invoking the protected function may
be scheduled off. Therefore, we do not allow any blocking
operations or system calls inside the protected function. As
mentioned earlier, we also disable paging for the processes
executing protected functions by using the mlockall()
system call with the MCL_FUTURE.

We detect whether a protected function has violated the
conditions of isolated execution by determining whether any
voluntary context switches occurred during the protected func-
tion’s execution. This usually indicates that either the protected
function yield the CPU manually or performed some blocking
operations.

Flushing shared resources. We modify the Linux scheduler
to check the taint of a core before scheduling a user process
on a processor core and to flush per-core resources if needed
as described in Section IV.

To flush the L1 and L2 caches, we iteratively read over
a segment of memory that is larger than the corresponding
cache sizes. We found this to be significantly more efficient
than using the WBINVD instruction, which we observed cost
as much as 300 microseconds in our tests. We flush the
L1 instruction cache by executing a large number of NOP
instructions.

Current implementations of Linux flush the TLB during
each context switch. Therefore, we do not need to separately
flush them. However, if Linux starts leveraging the PCID
feature of x86 processors in the future, the TLB would have
to be flushed explicitly. For flushing the BTB, we leveraged
a “branch slide” consisting of alternating conditional branch
and NOP instructions.

VI. EVALUATION

To show that our approach can be applied to protect a
wide variety of software, we have evaluated our solution in
three different settings and found that our solution successfully
prevents local and remote timing attacks in all of these settings.
We describe the settings in detail below.

Encryption algorithms implemented in high level interpreted
languages like Java. Traditionally, cryptographic algorithms
implemented in interpreted languages like Java have been
harder to protect from timing attacks than those implemented
in low level languages like C. Most interpreted languages
are compiled down to machine code on-the-fly by a VM
using Just-in-Time (JIT) code compilation techniques. The

JIT compiler often optimizes the code non-deterministically
to improve performance. This makes it extremely hard for
a programmer to reason about the transformations that are
required to make a sensitive function’s timing behavior secret-
independent. While developers writing low level code can
use features such as in-line assembly to carefully control the
machine code of their implementation, such low level control
is simply not possible in a higher level language.

We show that our techniques can take care of these issues.
We demonstrate that our defense can make the computation
time of Java implementations of cryptographic algorithms
independent of the secret key with minimal performance
overhead.

Cryptographic operations and SSL/TLS state machine. Im-
plementations of cryptographic primitives other than the pub-
lic/private key encryption or decryption routines may also
suffer from side channel attacks. For example, a cryptographic
hash algorithm like SHA-1 takes different amount of time
depending on the length of the input data. In fact, such timing
variations have been used as part of several existing attacks
against SSL/TLS protocols (e.g., Lucky 13). Also, the time
taken to perform the computation for implementing different
stages of the SSL/TLS state machine may also be dependent
on the secret key.

We find that our protection mechanism can protect cryp-
tographic primitives like hash functions as well as individual
stages of the SSL/TLS state machine from timing attacks while
incurring minimal overhead.

Sensitive data structures. Besides cryptographic algorithms,
timing channels also occur in the context of different data
structure operations like hash table lookups. For example, hash
table lookups may take different amount of time depending on
how many items are present in the bucket where the desired
item is located. It will take longer time to find items in buckets
with higher number of items than in the ones with less items.
This signal can be exploited by an attacker to cause denial of
service attacks [22]. We demonstrate that our technique can
prevent timing leaks using the associative arrays in C++ STL,
a popular hash table implementation.

Experiment setup. We perform all our experiments on a
machine with 2.3GHz Intel Xeon E5-2630 CPUs organized
in 2 sockets each containing 6 physical cores unless otherwise
specified. Each core has a 32KB L1 instruction cache, a 32KB
L1 data cache, and a 256KB L2 cache. Each socket has a
15MB L3 cache. The machine has a total of 64GB of RAM.

For our experiments, we use OpenSSL version 1.0.1l and
Java version BouncyCastle 1.52 (beta). The test machine runs
Linux kernel version 3.13.11.4 with our modifications as
discussed in Section V.

A. Security evaluation.

Preventing a simple timing attack. To determine the effective-
ness of our safe padding technique, we first test whether our
technique can protect against a large timing channel that can
distinguish between two different inputs of a simple function.
To make the attacker’s job easier, we craft a simple function
that has an easily observable timing channel—the function

0.00

0.05

0.10

0.15

0.20

0.25

0 20 40 60
Duration (ns)

F
re

q
u

e
n

cy

Input 0 1A. Unprotected

0.00

0.05

0.10

0.15

0.20

0.25

0 20 40 60
Duration (ns)

F
re

q
u

e
n

cy

B. With time padding but no randomized noise

0.00

0.05

0.10

0.15

0.20

2390 2400 2410
Duration (ns)

F
re

q
u

e
n

cy

C. Full protection (padding+randomized noise)

0.00

0.04

0.08

0.12

2390 2400 2410
Duration (ns)

F
re

q
u

e
n

cy

Fig. 4: Defeated distinguishing attack

executes a loop for 1 iteration if the input is 0 and 11 iterations
otherwise. We use the x86 loop instruction to implement
the loop and just a single nop instruction as the body of the
loop. We assume that the attacker calls the protected function
directly and measures the value of the timestamp counter
immediately before and after the call. The goal of the attacker
is to distinguish between two different inputs (0 and 1) by
monitoring the execution time of the function. Note that these
conditions are extremely favorable for an attacker.

We found that our defense completely defeats such a
distinguishing attack despite the highly favorable conditions
for the attacker. We also found that the timing randomization
step (described in Section IV-A) is critical for such protection
and a naive padding loop with any timing randomization step
indeed leaks information. Figure 4(A) shows the distributions
of observed runtimes of the protected function on inputs 0
and 1 with no defense applied. Figure 4(B) shows the runtime
distributions where padding is added to reach Tmax = 5000
cycles (≈ 2.17 µs) without the time randomization step. In
both cases, it can be seen that the observed timing distribu-
tions for the two different inputs are clearly distinguishable.
Figure 4(C) shows the same distributions when m = 5 rounds
of timing randomization are applied along with time padding.
In this case, we are no longer able to distinguish the timing
distributions.

We quantify the possibility of success for a distinguishing

−5

−4

−3

−2

−1

0

0 1 2 3 4 5
Rounds of noise

lo
g

1
0

(E
m

p
.
st

a
tis

tic
a

l d
is

ta
n

ce
)

Inputs

0 vs. 1

0 vs. 0

Fig. 5: The effect of multiple rounds of randomized noise addition
on the timing channel

attack in Figure 5 by plotting the variation of empirical
statistical distance between the observed distributions as the
amount of padding noise added is changed. The statistical
distance is computed using the following formula.

d(X,Y ) =
1

2


i∈Ω
|P [X = i]−P [Y = i]|

We measure the statistical distance over the set of observations
that are within the range of 50 cycles on either side of the me-
dian (this contains nearly all observations.) Each distribution
consist of around 600 million observations.

The dashed line in Figure 5 shows the statistical distance
between two different instances of the test function with 0 as
input. The solid line shows the statistical distance where one
instance has 0 as input and the other has 1. We observe that
the attack can be completely prevented if at least 2 rounds of
noise are used.

Preventing timing attack on RSA decryption We next evaluate
the effectiveness of our time padding approach to defeat
the timing attack by Brumley et al. [15] against unblinded
RSA implementations. Blinding is an algorithmic modification
to RSA that uses randomness to prevent timing attacks. To
isolate the impact of our specific defense, we apply our
defense to the RSA implementation in OpenSSL 1.0.1h with
such constant time defenses disabled. In order to do so, we
configure OpenSSL to disable blinding, use the non-constant
time exponentiation implementation, and use the non-word-
based Montgomery reduction implementation. We measure the
time of decrypting 256-byte messages with a random 2048-bit
key. We chose messages to have Montgomery representations
differing by multiples of 21016. Figure 6(A) shows the average
observed running time for such a decryption operation, which
is around 4.16 ms. The messages are displayed from left to
right in sorted order of how many Montgomery reductions
occur during the decryption. Each message was sampled
roughly 8,000 times and the samples were randomly split
into 4 sample sets. As observed by Brumley et al. [15], the
number of Montgomery reductions can be roughly determined

−1.0

−0.5

0.0

0.5

1.0

Message

D
u

ra
tio

n
(

n
s)

(+
~

4
.2

5
x

1
0

6
)

Trial 1 2 3 4A. Unprotected

−2000

−1000

0

1000

2000

Messages

D
u

ra
tio

n
(n

s)

(+
~

4
.1

6
x

1
0

6
)

B. Protected

−1.0

−0.5

0.0

0.5

1.0

Messages

D
u

ra
tio

n
(n

s)

(+
~

4
.2

5
x

1
0

6
)

Fig. 6: Protecting against timing attacks on unblinded RSA

from the running time of an unprotected RSA decryption. Such
information can be used to derive full length keys.

We then apply our defense to this decryption with Tmax
set to 9.68 × 106 cycles ≈ 4.21 ms. One timer interrupt
is guaranteed to occur during such an operation, as timer
interrupts occur at a rate of 250/s on our target machine. We
collect 30 million measurements and observe a multi-modal
padded distribution with four narrow, disjoint peaks corre-
sponding to the padding algorithm using different Text preempt
values for 1, 2, 3, and 4 interrupts respectively. The four
peaks represent, respectively, 94.0%,5.8%,0.6%, and 0.4% of
the samples. We did not observe that these probabilities vary
across different messages. Hence, in Figure 6(B), we show
the average observed time considering only observations from
within the first peak. Again, samples are split into 4 random
sample sets, each key is sampled around 700,000 times. We
observe no message-dependent signal.

Preventing cache attacks on AES encryption. We next
verify that our system protects against local cache attacks.
Specifically, we measured the effectiveness of our defense
against the PRIME+PROBE attack by Osvik et.al [35] on the
software implementation of AES encryption in OpenSSL. For
our tests, we apply the attack on only the first round of AES
instead of the full AES to make the conditions very favorable
to the attacker as subsequent rounds of AES add more noise to
the cache readings. In this attack, the attacker first primes the
cache by filling a selection of cache sets with the attacker’s
memory lines. Next, the attacker coerces the victim process
to perform an AES encryption on a chosen plaintext on the
same processor core. Finally, the attacker reloads the memory
lines it used to fill the cache sets prior to the encryption. This
allows the attacker to detect whether the reloaded lines were
still cached by monitoring timing or performance counters and
thus infer which memory lines were accessed during the AES
encryption operation.

On our test machine, the OpenSSL software AES imple-

A. Unprotected

0

5

10

15

0 10 20 30 0 10 20 30

Cache set

p
i

/
1

6

B. Protected

0

5

10

15

0 10 20 30 0 10 20 30

Cache set

p
i

/
1

6

Fig. 7: Protecting against cache attacks on software AES

mentation performs table lookups during the first round of
encryption that access one of 16 cache sets in each of 4 lookup
tables. The actual cache sets accessed during the operation are
determined by XORs of the top 4 bits of certain plaintext
bytes pi and certain key bytes ki. By repeatedly observing
cache accesses on chosen plaintexts where pi takes all possible
values of its top 4 bits, but where the rest of the plaintext is
randomized, the attacker observes cache line access patterns
revealing the top 4 bits of pi ⊕ ki, and hence the top 4 bits
of the key ki. This simple attack can be extended to learn the
entire AES key.

We use a hardware performance monitoring counter that
counts L2 cache misses as the probe measurement, and for
each measurement we subtract off the average measurement for
that cache set for all values of pi. Figure 7(A) and Figure 7(B)
show the probe measurements when performing this attack for
all values of the top 4 bits of p0 (left) and p5 (right) with
and without our protection scheme, respectively. Darker cells
indicate elevated measurements, and hence imply cache sets
that contain a line loaded by the attacker during the “prime”
phase that was evicted by the AES encryption. The secret key
k is randomly chosen, except that k0 = 0 and k5 = 80dec.
Without our solution, the cache set accesses show a pattern
revealing pi ⊕ ki which can be used to determine that the
top 4 bits of k0 and k5 are indeed 0 and 5, respectively. Our
solution flushes the L2 cache lazily before handing it over
to any untrusted process and thus ensures that no signal is
observed by the attacker as shown in Figure 7(B).

B. Performance evaluation

Performance costs of individual components. Table I shows
the individual cost of the different components of our defense.
Our total performance overhead is less than the total sum
of these components as we do not perform most of these
operations in the critical path. Note that retrieving the number
of times a process was interrupted or determining whether a
voluntary context switch occurred during a protected function’s

Component Cost (ns)
m = 5 time randomization step, WCET 710
Get interrupt counters 16
Detect context switch 4
Set and restore SCHED FIFO 2,650
Set and restore CPU affinity 1,235
Flush L1D+L2 cache 23,000
Flush BTB cache 7,000

TABLE I: Performance overheads of individual components of our
defense. WCET indicates worst-case execution time. Only costs listed
in the upper half of the table are incurred on each call to a protected
function.

execution is negligible due to our modifications to the Linux
kernel described in Section V.

Microbenchmarks: cryptographic operations in multiple lan-
guages. We perform a set of microbenchmarks that test
the impact of our solution on individual operations such as
RSA and ECDSA signing in the OpenSSL C library and in
the BouncyCastle Java library. In order to apply our defense
to BouncyCastle, we constructed JNI wrapper functions that
call the fixed time begin and fixed time end functions. Since
both libraries implement RSA blinding to defend against
timing attacks, we disable RSA blinding when applying our
defense.

The results of the microbenchmarks are shown in Table II.
Note that the delays experienced in any real applications will
be significantly less than these micro benchmarks as real
applications will also perform some I/O operations that will
amortize the performance overhead.

For OpenSSL, our solution adds between 3% (for RSA)
and 71% (for ECDSA) to the cost of computing a signature on
average. However, we offer significantly reduced tail latency
for RSA signatures. This behavior is caused by the fact that
OpenSSL regenerates the blinding factors every 32 calls to
the signing function to amortize the performance cost of
generating the blinding factors.

Focusing on the BouncyCastle results, our solution results
in a 2% decrease in cost for RSA signing and a 63% in-
crease in cost for ECDSA signing, compared to the stock
BouncyCastle implementation. We believe that this increase
in cost for ECDSA is justified by the increase in security,
as the stock BouncyCastle implementation does not defend
against local timing attacks. Furthermore, we believe that some
optimizations, such as configuring the Java VM to schedule
garbage collection outside of protected function executions,
could reduce this overhead.

Macrobenchmark: protecting the TLS state machine. We
applied our solution to protect the server-side implementation
of the TLS connection protocol in OpenSSL. The TLS protocol
is implemented as a state machine in OpenSSL, and this pre-
sented a challenge for applying our solution which is defined in
terms of protected functions. Additionally, reading and writing
to a socket is interleaved with cryptographic operations in the
specification of the TLS protocol, which conflicts with our
solution’s requirement that no blocking I/O may be performed
within a protected function.

RSA 2048-bit sign Mean (ms) 99% Tail
OpenSSL w/o blinding 1.45 1.45
Stock OpenSSL 1.50 2.18
OpenSSL + our solution 1.55 1.59
BouncyCastle w/o blinding 9.02 9.41
Stock BouncyCastle 9.80 10.20
BouncyCastle + our solution 9.63 9.82
ECDSA 256-bit sign Mean (ms) 99% Tail
Stock OpenSSL 0.07 0.08
OpenSSL + our solution 0.12 0.38
Stock BouncyCastle 0.22 0.25
BouncyCastle + our solution 0.36 0.48

TABLE II: Impact on performance of signing a 100 byte message
using SHA-256 with RSA or ECDSA for the OpenSSL and Boun-
cyCastle implementations. Measurements are in milliseconds. We
disable blinding when applying our defense to the RSA signature
operation. Bold text indicates a measurement where our defense
results in better performance than the stock implementation.

We addressed both challenges by generalizing the notion of
a protected function to that of a protected interval, which is an
interval of execution starting with a call to fixed time begin
and ending with fixed time end. We then split an execution
of the TLS protocol into protected intervals on boundaries
defined by transitions of the TLS state machine and on low-
level socket read and write operations. To achieve this, we
first inserted calls to fixed time begin and fixed time end at
the start and end of each state within the TLS state machine
implementation. Next, we modified the low-level socket read
and socket write OpenSSL wrapper functions to end the current
interval, communicate with the socket, and then start a new
interval. Thus divided, all cryptographic operations performed
inside the TLS implementation are within a protected interval.
Each interval is uniquely identifiable by the name of the
current TLS state concatenated with an integer incremented
every time a new interval is started within the same TLS state
(equivalently, the number of socket operations that occurred
so far during the state.)

The advantage of this strategy is that, unlike any prior
defenses, it protects the entire implementation of the TLS
state machine from any form of timing attack. However, such
protection schemes may incur additional overheads due to
protecting parts of the protocol that may not be vulnerable
to timing attacks because they do not work with secret data.

We evaluate the performance of the fully protected TLS
state machine as well as an implementation that only protects
the public key signing operation. The results are shown in Ta-
ble III. We observe an overhead of less than 5% on connection
latency even when protecting the full TLS protocol.

Protecting sensitive data structures. We measured the over-
head of applying our approach to protect the lookup operation
of the C++ STL unordered_map. For this experiment, we
populate the hash map with 1 million 64-bit integer keys and
values. We assume that the attacker cannot insert elements
in the hash map or cause collisions. The average cost of
performing a lookup of a key present in the map is 0.173µs
without any defense and 2.46µs with our defense applied.
Most of this overhead is caused by the fact that the worst-case
execution time of the lookup operation is significantly larger

Connection latency (RSA) Mean (ms) 99% Tail
Stock OpenSSL 5.26 6.82
Stock OpenSSL+ Our solution
(sign only)

5.33 6.53

Stock OpenSSL+ Our solution 5.52 6.74
Connection latency (ECDSA) Mean (ms) 99% Tail
Stock OpenSSL 4.53 6.08
Stock OpenSSL+ Our solution
(sign only)

4.64 6.18

Stock OpenSSL+ Our solution 4.75 6.36

TABLE III: The impact on TLS v1.2 connection latency when apply-
ing our defense to the OpenSSL server-side TLS implementation.
We evaluate the cases where the the server uses an RSA 2048-
bit or ECDSA 256-bit signing key with SHA-256 as the digest
function. Latency given in milliseconds and measures the end-to-end
connection time. The client uses the unmodified OpenSSL library
attempts. We evaluate our defense when only protecting the signing
operation and when protecting all server-side routines performed as
part of the TLS connection protocol that use cryptography. Even when
the full TLS protocol is protected, our approach adds an overhead of
less than 5% to average connection latency. Bold text indicates a
measurement where our defense results in better performance than
the stock implementation.

than the average-case. the profiled worst-case execution time of
the lookup when interrupts do not occur is 1.32µs at κ = 10−5.
Thus, any timing channel defense will cause the lookup to
take at least 1.32µs. The worst-case execution estimate of the
lookup operation increases to 13.3µs when interrupt cases are
not excluded, hence our scheme benefits significantly from
adapting to interrupts during padding for this example. Another
major part of the overhead of our solution (0.710µs) comes
from the randomization step to ensure safe padding . As we
described earlier in Section VI-A, the randomization step is
crucial to ensure that there is no timing leakage.

Hardware portability. Our solution is not specific to any
particular hardware. It will work on any hardware that supports
standard cache hierarchy and where page coloring can be im-
plemented. To test the portability of our solution, we executed
some of the benchmarks mentioned in Sections VI-A and VI-B
on a 2.93 GHz Intel Xeon X5670 CPU. We confirmed that
our solution successfully protects against the local and remote
timing attacks on that platform too. The relative performance
overheads were similar to the ones reported above.

VII. LIMITATIONS

No system calls inside protected functions. Our current
prototype does not support protected functions that invoke
system calls. A system call can inadvertently leak information
to an attacker by leaving state in shared kernel data structures,
which an attacker might indirectly observe by invoking the
same system call and timing its duration. Alternatively, a
system call might access regions of the L3 cache that can
be snooped by an attacker process.

The lack of system call support turned out to be not a big
issue in practice as our experiments so far indicate that system
calls are rarely used in functions dealing with sensitive data
(e.g., cryptographic operations). However, if needed in future,
one way of supporting system calls inside protected functions

while still avoiding this leakage is to apply our solution to the
kernel itself. For example, we can pad any system calls that
modify some shared kernel data structures to their worst case
execution times.

Indirect timing variations in unprotected code. Our ap-
proach does not currently defend against timing variations in
the execution of non-sensitive code segments that might get
indirectly affected by a protected function’s execution. For
example, consider the case where a non-sensitive function
from a process gets scheduled on a processor core immediately
after another process from the same user finishes executing a
protected function. In such a case, our solution will not flush
the state of per-core resources like L1 cache as both these
processes belong to the same user. However, if such remnant
cache state affects the timing of the non-sensitive function, an
attacker may be able to observe these variations and infer some
information about the protected function.

Note that currently there are no known attacks that could
exploit this kind of leakage. A conservative approach that
prevents such leakages is to flush all per-cpu resources at the
end of each protected function. This will, of course, result
in higher performance overheads. The costs associated with
cleansing different types of per-cpu resources are summarized
in Table I.

Leakage due to fault injection. If an attacker can cause
a process to crash in the middle of a protected function’s
execution, the attacker can potentially learn secret information.
For example, consider a protected function that first performs
a sensitive operation and then parses some input from the user.
An attacker can learn the duration of the sensitive operation
by providing a bad input to the parser that makes it crash and
measuring how long it takes the victim process to crash.

Our solution, in its current form, does not protect against
such attacks. However, this is not a fundamental limitation.
One simple way of overcoming these attacks is to modify
the OS to apply the time padding for a protected function
even after it has crashed as part of the OS’s cleanup handler.
This can be implemented by modifying the OS to keep track
of all processes that are executing protected functions at any
given point of time and their respective padding parameters.
If any protected function crashes, the OS cleanup handler for
the corresponding process can apply the desired amount of
padding.

VIII. RELATED WORK

A. Defenses against remote timing attacks

The remote timing attacks exploit the input-dependent exe-
cution times of cryptographic operations. There are three main
approaches to make cryptographic operations’ execution times
independent of their inputs: static transformation, application-
specific changes, and dynamic padding.

Application-specific changes. One conceptually simple way
to defend an application against timing attacks is to modify its
sensitive operations such that their timing behavior is not key-
dependent. For example, AES [10, 27, 30] implementations
can be modified to ensure that their execution times are
key-independent. Note that, since the cache behavior impacts

running time, achieving secret-independent timing usually re-
quires rewriting the operation so that its memory access pattern
is also independent of secrets. Such modifications are applica-
tion specific, hard to design, and very brittle. By contrast, our
solution is completely independent of the application and the
programming language.

Static transformation. An alternative approach to prevent
remote attacks is to use static transformations on the imple-
mentation of the cryptographic operation to make it constant
time. One can use a static analyzer to find the longest possible
path through the cryptographic operation and insert padding
instructions that have no side-effects (like NOP) along other
paths so that they take the same amount of time as the longest
path [17, 20]. While this approach is generic and can be
applied to any sensitive operation, it has several drawbacks. In
modern architectures like x86, the execution time of several
instructions (e.g., the integer divide instruction and multiple
floating-point instructions) depend the value of the input of
these instructions. This makes it extremely hard and time
consuming to statically estimate the execution time of these
instructions. Moreover, it is very hard to statically predict the
changes in the execution time due to internal cache collisions
in the implementation of the cryptographic operation. To avoid
such cases, in our solution, we use dynamic offline profiling
to estimate the worst-case runtime of a protected function.
However, such dynamic techniques suffer from incompleteness
i.e. they might miss worst-case execution times triggered by
pathological inputs.

Dynamic padding. Dynamic padding techniques add a vari-
able amount of padding to a sensitive computation that depends
on the observed execution time of the computation in order
to mitigate the timing side-channel. Several prior works [6,
18, 24, 31, 47] have presented ways to pad the execution of a
black-box computation to certain predetermined thresholds and
obtain bounded information leakage. Zhang et al. designed a
new programming language that, when used to write sensitive
operations, can enforce limits on the timing information leak-
age [48]. The major drawback of existing dynamic padding
schemes is that they incur large performance overhead. This
results from the fact that their estimations of the worst-case
execution time tend to be overly pessimistic as it depends on
several external parameters like OS scheduling, cache behavior
of the simultaneously running programs, etc. For example,
Zhang et al. [47] set the worst-case execution time to be 300
seconds for protecting a Wiki server. Such overly pessimistic
estimates increase the amount of required padding and thus
results in significant performance overheads (90 − 400% in
macro-benchmarks [47]). Unlike existing dynamic padding
schemes, our solution incurs minimal performance overhead
and protects against both local and remote timing attacks.

B. Defenses against local attacks

Local attackers can also perform timing attacks, hence
some of the defenses provided in the prior section may also
be used to defend against some local attacks. However, local
attackers also have access to shared hardware resources that
contain information related to the target sensitive operation.
The local attackers also have access to fine-grained timers.

A common local attack vector is to probe a shared hardware

resource, and then, using the fine-grained timer, measure how
long the probe took to run. Most of the proposed defenses to
such attacks try to either remove access to fine-grained timers
or isolate access to the shared hardware resources. Some of
these defenses also try to minimize information leakage by
obfuscating the sensitive operation’s hardware access patterns.
We describe these approaches in detail below.

Removing fine-grained timers. Several prior projects have
evaluated removing or modifying time measurements taken on
the target machine [33, 34, 42]. Such solutions are often quite
effective at preventing a large number of local side channel
attacks as the underlying states of most shared hardware
resources can only be read by accurately measuring the time
taken to perform certain operations (e.g., read a cache line).

However, removing access to wall clock time is not suffi-
cient for protecting against all local attackers. For example, a
local attacker executing multiple probe threads can infer time
measurements by observing the scheduling behavior of the
threads. Custom scheduling schemes (e.g., instruction-based
scheduling) can eliminate such an attack [38] but implementing
these defenses require major changes to the OS scheduler. In
contrast, our solution only requires minor changes to the OS
scheduler and protects against both local and remote attackers.

Preventing sharing of hardware state across processes. Many
proposed defenses against local attackers prevent an attacker
from observing state changes to shared hardware resources
caused by a victim process. We divide the proposed defenses
into five categories and describe them next.

Resource partitioning. Partitioning shared hardware resources
can defeat local attackers, as they cannot access the same
partition of the resource as a victim. Kim et al. [28] present
an efficient management scheme for preventing local timing
attacks across virtual machines (VMs). Their technique locks
memory regions accessed by sensitive functions into reserved
portions of the L3 cache. This scheme can be more efficient
than page coloring. Such protection schemes are comple-
mentary to our technique. For example, our solution can be
modified to use such a mechanism instead of page coloring to
dynamically partition the L3 cache.

Some of the other resource partitioning schemes (e.g.,
Ristenpart et al. [37]) suggest allocating dedicated hardware
to each virtual machine instance to prevent cross-VM attacks.
However, such schemes are wasteful of hardware resources as
they decrease the amount of resources available to concurrent
processes. By contrast, our solution utilizes the shared hard-
ware resources efficiently as they are only isolated during the
execution of the protected functions. The time a process spends
executing protected functions is usually much smaller than the
time it spends in non-sensitive computations.

Limiting concurrent access. If gang scheduling [28] is used
or hyperthreading is disabled, an attacker can only observe
per-CPU resources when it has preempted a victim. Hence,
reducing the frequency of preemptions reduces the feasibility
of cache-attacks on per-CPU caches. Varadarajan et al. [41]
propose using minimum runtime guarantees to ensure that a
VM is not preempted too frequently. However, as noted in [41],
such a scheme is very hard to implement in a OS scheduler

as, unlike a hypervisor scheduler, an OS scheduler must deal
with a unbounded number of processes.

Custom hardware. Custom hardware can be used to obfuscate
and randomize the victim process’s usage of the hardware. For
example, Wang et al. [43, 44] proposed new ways of designing
caches that ensures that no information about cache usage is
shared across different processes. However such schemes have
limited practical usage as they, by design, cannot be deployed
on off-the-shelf commodity hardware.

Flushing state. Another class of defenses ensure that the state
of any per-CPU hardware resources are cleared before trans-
ferring them from one process to another. Düppel, by Zhang
et al. [50], flushes per-CPU L1 and (optionally) L2 caches
periodically in a multi-tenant VM setting. Their solution also
requires the hyperthreading to be disabled. They report around
7% overheads on regular workloads. In essence, this scheme
is similar to our solution’s technique of flushing per-CPU
resources in the OS scheduler. However, unlike Düppel, we
flush the state lazily only when a context switch to a different
user process than the one executing a protected operation
occurs. Also, Düppel only protects against local cache attacks.
We protect against both local and remote timing and cache
attacks while still incurring less overhead than Düppel.

Application transformations. Sensitive operations like sensi-
tive computations in different programs can also be modified
to exhibit either secret-independent or obfuscated hardware
access patterns. If the access to the hardware is independent
of secrets, then an attacker cannot use any of the state leaked
through shared hardware to learn anything meaningful about
the sensitive operations. Several prior projects have shown
how to modify AES implementations to obfuscate their cache
access patterns [9, 10, 13, 35, 40]. Similarly, recent versions of
OpenSSL use a specifically modified implementation of RSA
that ensures secret-independent cache accesses. Some of these
transformations can also be applied dynamically. For example,
Crane et al. [21] implement a system that dynamically applies
cache-access obfuscating transformations to an application at
runtime.

However, these transformations are specific to particular
cryptographic operations and are very hard to implement and
maintain correctly. For example, 924 lines of assembly code
had to be added to OpenSSL to implement make the RSA
implementation’s cache accesses secret-independent.

IX. CONCLUSION

We presented a low-overhead, cross-architecture defense
that protects applications against both local and remote timing
attacks with minimal application code changes. Our exper-
iments and evaluation also show that our defense works
across different applications written in different programming
languages.

Our solution defends against both local and remote attacks
by using a combination of two main techniques: (i) a time
padding scheme that only takes secret-dependent time vari-
ations into account, and (ii) preventing information leakage
via shared resources such as the cache and branch prediction
buffers. We demonstrated that applying small time pads ac-
curately is non-trivial because the timing loop itself may leak

information. We developed a method by which small time pads
can be applied securely. We hope that our work will motivate
application developers to leverage some of our techniques to
protect their applications from a wide variety of timing attacks.
We also expect that the underlying principles of our solution
will be useful in future work protecting against other forms of
side channel attacks.

ACKNOWLEDGMENTS

This work was supported by NSF, DARPA, ONR, and a
Google PhD Fellowship to Suman Jana. Opinions, findings and
conclusions or recommendations expressed in this material are
those of the author(s) and do not necessarily reflect the views
of DARPA.

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  • I Introduction
  • II Known timing attacks
  • III Threat Model
  • IV Our Solution
    • IV-A Time padding
    • IV-B Preventing leakage through shared resources
  • V Implementation
    • V-A Programming API
    • V-B Time padding
    • V-C Prevent leakage through shared resources
  • VI Evaluation
    • VI-A Security evaluation.
    • VI-B Performance evaluation
  • VII Limitations
  • VIII Related work
    • VIII-A Defenses against remote timing attacks
    • VIII-B Defenses against local attacks
  • IX Conclusion

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