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title: "ASLR on the Line: Practical Cache Attacks on the MMU"
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    title: "ASLR on the Line: Practical Cache Attacks on the MMU"
    author: Ben Gras, Kaveh Razavi, Erik Bosman, Herbert Bos, Cristiano Giuffrida
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  - Kaveh Razavi
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  - Cristiano Giuffrida
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# ASLR on the Line: Practical Cache Attacks on the MMU

**ASLR on the Line: Practical Cache Attacks on the MMU** - Ben Gras, Kaveh Razavi, Erik Bosman, Herbert Bos, Cristiano Giuffrida, NDSS Symposium.

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# ASLR on the Line: Practical Cache Attacks on the MMU

--- page 1 ---

ASLR on the Line: Practical Cache Attacks on the MMU
Ben Gras

Kaveh Razavi

Erik Bosman Herbert Bos Cristiano Giuffrida
Vrije Universiteit Amsterdam
f
beng, kaveh, ejbosman, herbertb, giuffrida
g
@cs.vu.nl

Equal contribution joint rst authors
Abstract
—Address space layout randomization (ASLR) is an
important rst line of defense against memory corruption attacks
and a building block for many modern countermeasures. Existing
attacks against ASLR rely on software vulnerabilities and/or on
repeated (and detectable) memory probing.
In this paper, we show that neither is a hard requirement
and that ASLR is
fundamentally
insecure on modern cache-
based architectures, making ASLR and caching conicting
requirements (ASLR

Cache, or simply AnC). To support
this claim, we describe a new
EVICT+TIME
cache attack
on the virtual address translation performed by the memory
management unit (MMU) of modern processors. Our AnC attack
relies on the property that the MMU's page-table walks result
in caching page-table pages in the shared last-level cache (LLC).
As a result, an attacker can derandomize virtual addresses of a
victim's
code and data
by locating the cache lines that store the
page-table entries used for address translation.
Relying only on basic memory accesses allows AnC to be
implemented in JavaScript without any specic instructions or
software features. We show our JavaScript implementation can
break code and heap ASLR in two major browsers running on
the latest Linux operating system with 28 bits of entropy in 150
seconds. We further verify that the AnC attack is applicable to
every modern architecture that we tried, including Intel, ARM
and AMD. Mitigating this attack without naively disabling caches
is hard, since it targets the low-level operations of the MMU.
We conclude that ASLR is fundamentally awed in sandboxed
environments such as JavaScript and future defenses should not
rely on randomized virtual addresses as a building block.
I. I
NTRODUCTION
Address-space layout randomization (ASLR) is the rst
line of defense against memory-related security vulnerabilities
in today's modern software. ASLR selects random locations
in the large virtual address space of a protected process for
placing code or data. This simple defense mechanism forces
attackers to rely on secondary software vulnerabilities (e.g.,
arbitrary memory reads) to directly leak pointers [16], [56] or
ad-hoc mechanisms to bruteforce the randomized locations [5],
[6], [17], [19], [23], [47], [54].
Finding secondary information leak vulnerabilities raises
the effort on an attacker's side for exploitation [22]. Also
bruteforcing, if at all possible [16], [59], requires repeat-
edly generating anomalous events (e.g., crashes [5], [17],
[54], exceptions [19], [23], or huge allocations [47]) that are
easy to detect or prevent. For instance, for some attacks [6]
disabling non-fundamental memory management features is
enough [62]. Consequently, even if ASLR does not stop the
more advanced attackers, in the eyes of many, it still serves
as a good rst line of defense for protecting the users and
as a pivotal building block in more advanced defenses [9],
[15], [36], [42], [52]. In this paper, we challenge this belief by
systematically derandomizing ASLR through a side-channel
attack on the memory management unit (MMU) of processors
that we call ASLR

Cache (or simply AnC).
Previous work has shown that ASLR breaks down in
the presence of specic weaknesses and (sometimes arcane)
features in software. For instance, attackers may derandomize
ASLR if the application is vulnerable to thread spraying [23],
if the system turns on memory overcommit and exposes
allocation oracles [47], if the application allows for crash
tolerant/resistant memory probing [5], [17], [19], [54], or if the
underlying operating system uses deduplication to merge data
pages crafted by the attacker with pages containing sensitive
system data [6]. While all these conditions hold for
some
applications, none of them are universal and they can be
mitigated in software.
In this paper, we show that the problem is much more
serious and that ASLR is
fundamentally insecure
on modern
cache-based architectures. Specically, we show that it is
possible to derandomize ASLR completely from JavaScript,
without resorting to esoteric operating system or application
features. Unlike all previous approaches, we do not abuse
weaknesses in the software (that are relatively easy to x).
Instead, our attack builds on hardware behavior that is central
to efcient code execution: the fast translation of virtual to
physical addresses in the MMU by means of page tables.
As a result, all xes to our attacks (e.g., naively disabling
caching) are likely too costly in performance to be practical.
To our knowledge, this is the rst attack that side-channels
the MMU and also the very rst cache attack that targets a
victim hardware rather than software component.
High level overview of the attack
Whenever a process wants
to access a virtual address, be it data or code, the MMU per-
forms a translation to nd the corresponding physical address
in main memory. The translation lookaside buffer (TLB) in
each CPU core stores most of the recent translations in order
to speed up the memory access. However, whenever a TLB
miss occurs, the MMU needs to walk the page tables (PTs) of
the process (also stored in main memory) to perform the trans-
Permission to freely reproduce all or part of this paper for noncommercial
purposes is granted provided that copies bear this notice and the full citation
on the rst page. Reproduction for commercial purposes is strictly prohibited
without the prior written consent of the Internet Society, the rst-named author
(for reproduction of an entire paper only), and the author's employer if the
paper was prepared within the scope of employment.
NDSS '17, 26 February - 1 March 2017, San Diego, CA, USA
Copyright 2017 Internet Society, ISBN 1-1891562-46-0
http://dx.doi.org/10.14722/ndss.2017.23271

--- page 2 ---

lation. To improve the performance of the MMU walk (i.e., a
TLB miss), the PTs are cached in the fast data caches very
much like the process data is cached for faster access [4], [28].
Relying on ASLR as a security mechanism means that
the PTs now store security-sensitive secrets: the offset of a
PT entry in a PT page at each PT level encodes part of
the secret virtual address. To the best of our knowledge, the
implications of sharing the CPU data caches between the secret
PT pages and untrusted code (e.g., JavaScript code) has never
been previously explored.
By executing specially crafted memory access patterns on
a commodity Intel processor, we are able to infer which cache
sets have been accessed after a targeted MMU PT walk when
dereferencing a data pointer or executing a piece of code. As
only certain addresses map to a specic cache set, knowing
the cache sets allows us to identify the offsets of the target PT
entries at each PT level, hence derandomizing ASLR.
Contributions
Summarizing, we make the following contri-
butions:
1) We design and implement AnC, the rst cache
side-channel attack against a hardware component
(i.e., the processor's MMU), which allows malicious
JavaScript code to derandomize the layout of the
browser's address space, solely by accessing memory.
Since AnC does not rely on any specic instruction or
software feature, it cannot be easily mitigated without
naively disabling CPU caches.
2) To implement AnC, we needed to implement better
synthetic timers than the one provided by the current
browsers. Our timers are practical and can tell the
difference between a cached and an uncached mem-
ory access. On top of AnC, they make previous cache
attacks (e.g., [48]) possible again.
3) While AnC fundamentally breaks ASLR, we fur-
ther show, counter-intuitively perhaps, that memory
allocation patterns and security countermeasures in
current browsers, such as randomizing the location
of the browser heaps on every allocation, make AnC
attacks more effective.
4) We evaluated end-to-end attacks with AnC on two
major browsers running on Linux. AnC runs in tens
of seconds and successfully derandomizes code and
heap pointers, signicantly reducing an attacker's
efforts to exploit a given vulnerability.
Outline
After presenting the threat model in Section II, we
explain the details of address translation in Section III. In
that section, we also summarize the main challenges and
how we address them. Next, Sections IV—VI discuss our
solutions for each of the challenges in detail. In Section VII,
we evaluate AnC against Chrome and Firefox, running on the
latest Linux operating system. We show that AnC successfully
derandomizes ASLR of the heap in both browsers and ASLR
of the JIT code in Firefox while being much faster and
less demanding in terms of requirements than state-of-the-
art derandomization attacks. We discuss the impact of AnC
on browser-based exploitation and on security defenses that
rely on information hiding in the address space or leakage-
resilient code randomization in Section VIII. We then propose
mitigations to limit (but not eliminate) the impact of the attack
in Section IX and highlight the related work in Section X
before concluding in Section XI.
II. T
HREAT
M
ODEL
We assume the attacker can execute JavaScript code in
the victim's browser, either by luring the victim into visiting
a malicious website or by compromising a trusted website.
Assuming all the common defenses (e.g., DEP) are enabled in
the browser, the attacker aims to escape the JavaScript sand-
box via a memory corruption vulnerability. To successfully
compromise the JavaScript sandbox, we assume the attacker
needs to rst break ASLR and derandomize the location
of some code and/or data pointers in the address space—
a common attack model against modern defenses [53]. For
this purpose, we assume the attacker cannot rely on ad-hoc
disclosure vulnerabilities [16], [56] or special application/OS
behavior [5], [6], [17], [19], [23], [47], [54]. While we focus
on a JavaScript sandbox, the same principles apply to other
sandboxing environments such as Google's Native Client [65].
III. B
ACKGROUND AND
A
PPROACH
In this section, we discuss necessary details of the memory
architecture in modern processors. Our description will focus
on recent Intel processors due to their prevalence, but other
processors use similar designs [4] and are equally vulnerable
as we will show in our evaluation in Section VII.
A. Virtual Address Translation
Currently, the virtual address-space of a 64 bit process is
256 TB on x8664 processors, whereas the physical memory
backing it is often much smaller and may range from a few
KBs to a few GBs in common settings. To translate a virtual
address in the large virtual address-space to its corresponding
physical address in the smaller physical address-space, the
MMU uses the PT data structure.
The PT is a uni-directed tree where parts of the virtual
address select the outgoing edge at each level. Hence, each
virtual address uniquely selects a path between the root of the
tree to the leaf where the target physical address is stored.
As the current x8664 architecture uses only the lower 48 bits
for virtual addressing, the total address space is 256 TB. Since
a PT maintains a translation at the granularity of a memory
page (4096 bytes), the lower 12 bits of a virtual page and its
corresponding physical page are always the same. The other
36 bits select the path in the PT tree. The PT tree has four
levels of page tables, where each PT is itself a page that stores
512 PT entries (PTEs). This means that at each PT level, 9 of
the aforementioned 36 bits decide the offset of the PTE within
the PT page.
Figure 1 shows how the MMU uses the PT for translating
an example virtual address, 0x644b321f4000. On the x86
architecture, the CPU's
CR3
control register points to the
highest level of the page table hierarchy, known as level 4
or PTL4. The top 9 bits of the virtual address index into this
single PTL4 page, in this case selecting PTE 200. This PTE
has a reference to the level 3 page (i.e., PTL3), which the
next 9 bits of the virtual address index to nd the target PT
entry (this time at offset 300). Repeating the same operation
2

--- page 3 ---

Fig. 1. MMU's page table walk to translate 0x644b321f4000 to its
corresponding memory page on the x8664 architecture.
for PT pages at level 2 and 1, the MMU can then nd the
corresponding physical page for 0x644b321f4000 at the PT
entry in the level 1 page.
Note that each PTE will be in a cache line, as shown by
different colors and patterns in Figure 1. Each PTE on x8664
is eight bytes, hence, each 64 byte cache line stores eight PTE.
We will discuss how we can use this information for derandom-
izing ASLR of a given virtual address in Section III-D after
looking into the memory organization and cache architecture
of recent Intel x8664 processors.
B. Memory Organization
Recent commodity processors contain a complex memory
hierarchy involving multiple levels of caches in order to speed
up the processor's access to main memory. Figure 2 shows how
the MMU uses this memory hierarchy during virtual to physi-
cal address translation in a recent Intel Core microarchitecture.
Loads and stores as well as instruction fetches on virtual
addresses are issued from the core that is executing a process.
The MMU performs the translation from the virtual address to
the physical address using the TLB before accessing the data
or the instruction since the caches that store the data are tagged
with physical addresses (i.e., physically-tagged caches). If the
virtual address is in the TLB, the load/store or the instruction
fetch can proceed. If the virtual address is not in the TLB, the
MMU needs to walk the PT as we discussed in Section III-A
and ll in the TLB. The TLB may include translation caches
for different PT levels (e.g., paging-structure caches on Intel
described in Section 4.10.3 of [29]). As an example, if TLB
includes a translation cache for PTL2, then the MMU only
needs to walk PTL1 to nd the target physical address.
During the PT walk, the MMU reads PT pages at each
PT level using their physical addresses. The MMU uses the
same path as the core for loading data to load the PTEs
during translation. As a result, after a PT walk, the cache
lines that store the PTE at each PT level are available in the
L1 data cache (i.e., L1D). We now briey discuss the cache
architecture.
C. Cache Architecture
In the Intel Core microarchitecture, there are three levels
of CPU caches
1
. The caches that are closer to the CPU are
smaller and faster whereas the caches further away are slower
but can store a larger amount of data. There are two caches1
The mobile version of the Skylake processors has a level 4 cache too.
Fig. 2. Memory organization in a recent Intel processor.
at the rst level, L1D and L1I, to cache data and instructions,
respectively. The cache at the second level, L2, is a unied
cache for both data and instructions. L1 and L2 are private to
each core, but all cores share L3. An important property of
these caches is their inclusivity. L2 is exclusive of L1, that is,
the data present in L1 is not necessarily present in L2. L3,
however, is inclusive of L1 and L2, meaning that if data is
present in L1 or L2, it also has to be present in L3. We later
exploit this property to ensure that a certain memory location
is not cached at any level by making sure that it is not present
in L3. We now discuss the internal organization of these CPU
caches.
To adhere to the principle of locality while avoiding expen-
sive logic circuits, current commodity processors partition the
caches at each level. Each partition, often referred to as a
cache
set
, can store only a subset of physical memory. Depending
on the cache architecture, the physical or virtual address of a
memory location decides its cache set. We often associate a
cache set with
wayness
. An
n-way
set-associative cache can
store
n
items in each cache set. A replacement policy then
decides which of the
n
items to replace in case of a miss in
a cache set. For example, the L2 cache on an Intel Skylake
processor is 256 KB and 4-way set-associative with a cache
line size of 64 B [28]. This means that there are 1024 cache
sets (256 KB/(4-way

64 B)) and bits 6 to 16 of a physical
address decide its corresponding cache set (the lower 6 bits
decide the offset within a cache line).
In the Intel Core microarchitecture, all the cores of the
processor share the LLC, but the microarchitecture partitions
it in so-called slices, one for each core, where each core has
faster access to its own slice than to the others. In contrast to
L1 and L2 where the lower bits of a physical address decide
its corresponding cache set, there is a complex addressing
function (based on an XOR scheme) that decides the slice for
each physical memory address [27], [44]. This means that each
slice gets different cache sets. For example, a 4-core Skylake
i7-6700K processor has an 8 MB 16-way set associative LLC
with 4 slices each with 2048 cache sets. We now show how PT
pages are cached and how we can evict them from the LLC.
D. Derandomizing ASLR
As discussed earlier, any memory access that incurs a TLB
miss requires a PT walk. A PT walk reads four PTEs from
main memory and stores them in four different cache lines in
L1D if they are not there already. Knowing the offset of these
3

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--- page 6 ---

cache lines within a page already derandomizes six bits out of
nine bits of the virtual address at each PT level. The last three
bits will still not be known because the offset of the PTE within
the cache line is not known. We hence need to answer three
questions in order to derandomize ASLR: (1) which cache
lines are loaded from memory during the PT walk, (2) which
page offsets do these cache lines belong to, and (3) what are
the offsets of the target PTEs in these cache lines?
1) Identifying the cache lines that host the PTEs:
Since
the LLC is inclusive of L1D, if the four PTEs cache lines are
in L1D, they will also be in the LLC and if they are not in
the LLC, they will also not be in L1D. This is an important
property that we exploit for implementing AnC: rather than
requiring a timer that can tell the difference between L1D and
LLC (assuming no L2 entry), we only require one that can
tell the difference between L1D and memory by evicting the
target cache line from the LLC rather than from L1D.
The PTE cache lines could land in up to four different
cache sets. While we cannot directly identify the cache lines
that host the PTE, by monitoring (or controlling) the state of
various cache sets at the LLC, we can detect MMU activity due
to a PT walk at the affected cache sets. While the knowledge
of MMU activity on cache sets is coarser than on cache lines,
it is still enough to identify the offset of the PTE cache lines
within a page as we describe next.
2) Identifying page offsets of the cache lines:
Oren et
al. [48] realized that given two different (physical) memory
pages, if their rst cache lines (i.e., rst 64 B) belong to the
same cache set, then their other 63 cache lines share (different)
cache sets as well. This is due to the fact that for the rst cache
lines to be in the same cache set, all the bits of the physical
addresses of both pages that decide the cache set and the slice
have to be the same and an offset within both memory pages
will share the lower 12 bits. Given, for example, 8192 unique
cache sets, this means that there are 128 (8192/64) unique page
colors
in terms of the cache sets they cover.
This simple fact has an interesting implication for our
attack. Given an identied cache set with PT activity, we can
directly determine its page color, and more importantly, the
offset of the cache line that hosts the PT entry.
3) Identifying cache line offsets of the PT entries:
At this
stage, we have identied the cache sets for PTEs at each PT
level. To completely derandomize ASLR for a given virtual
address, we still need to identify the PTE offset within a cache
line (inside the identied cache set), as well as mapping each
identied cache set to a PT level.
We achieve both goals via accessing pages that are
x
bytes
apart from our target virtual address
v
. For example, the pages
that are 4 KB, 8 KB, ..., 32 KB away from
v
, are 1 to 8 PTEs
away from
v
at PTL1 and if we access them, we are ensured
to see a change in one of the four cache sets that show MMU
activity (i.e., the new cache set will directly follow the previous
cache set). The moving cache set, hence, uniquely identies
as the one that is hosting the PT entry for PTL1, and the point
when the change in cache set happens uniquely identies the
PT entry offset of
v
within its cache line, derandomizing the
unknown 3 least signicant bits in PTL1. We can apply the
same principle for nding the PT entry offsets at other PT
levels. We call this technique
sliding
and discuss it further in
Section V-E.
E. ASLR on Modern Systems
Mapped virtual areas for position-independent executables
in modern Linux systems exhibit 28 bit of ASLR entropy. This
means that the PTL1, PTL2 and PTL3 fully contribute to
creating 27 bits of entropy, but only the last bit of the PTE
offset in PTL4 contributes to the ASLR entropy. Nevertheless,
if we want to identify this last bit, since it falls into the lowest
three bits of the PTE offset (i.e., within a cache line), we
require a crossing cache set at PTL4. Each PTE at PTL4
maps 512 GB of virtual address-space, and hence, we need
a virtual mapping that crosses a 4 TB mark in order for a
cache set change to happen at PTL4. Note that a cache set
change in PTL4 results in cache sets changing in the other
levels as well. We will describe how we can achieve this
by exploiting the behavior of memory allocators in various
browsers in Section VI.
Note that the entropy of ASLR in Linux is higher than
other popular operating systems such as Windows 10 [45], [63]
which provides only 24 bits of entropy for the heap and 17–
19 bits of entropy for executables. This means that on Windows
10, PTL4 does not contribute to ASLR for the heap area.
Since each entry in PTL3 covers 1 GB of memory, a mapping
that crosses an 8 GB will result in cache set change at PTL3,
resulting in derandomization of ASLR. The lower executable
entropy means that it is possible to derandomize executable
locations on Windows 10 when crossing only the two lower
level (i.e., with 16 MB). In this paper we focus on the much
harder case of Linux which provides the highest entropy for
ASLR.
F. Summary of Challenges and Approach
We have discussed the memory hierarchy on modern
x8664 processors and the way in which an attacker can mon-
itor MMU activity to deplete ASLR's entropy. The remainder
of this paper revolves around the three main challenges that
we need to overcome to implement a successful attack:
C1 Distinguishing between a memory access and a cache
access when performed by the MMU in modern
browsers. To combat timing attacks from a sandboxed
JavaScript code [6], [48], browsers have decreased the
precision of the accessible timers in order to make it
difcult, if not impossible, for the attackers to observe
the time it takes to perform a certain action.
C2 Observing the effect of MMU's PT walk on the state
of the data cache. Recent work [48] shows that it is
possible to build cache eviction sets from JavaScript in
order to bring the last-level cache (LLC) to a known
state for a well-known
PRIME+PROBE
attack [39],
[49]. In a typical
PRIME+PROBE
attack, the victim
is a process running on a core, whereas in our attack
the victim is the MMU with a different behavior.
C3 Distinguishing between multiple PT entries that are
stored in the same cache line and uniquely identifying
PT entries that belong to different PT levels. On e.g.,
x8664 each PT entry is 8 bytes, hence, each 64-byte
4

--- page 7 ---

Fig. 3. Measuring the quality of the low-precision
performance.now()
in Chrome and Firefox.
cache line can store 8 PT entries (i.e., the 3 lower
bits of the PT entry's offset is not known). Therefore,
to uniquely identify the location of a target PT entry
within a PT page, we require the ability to access the
virtual addresses that correspond to the neighboring
PT entries in order to observe a cache line change.
Further, in order to derandomize ASLR, we need PT
entries to cross cache line at each PT level.
To address C1, we have created a new synthetic timer in
JavaScript to detect the difference between a memory and
a cache access. We exploit the fact that the available timer,
although coarse, is precise and allows us to use the CPU core
as a counter to measure how long each operation takes. We
elaborate on our design and its implications on browser-based
timing attacks in Section IV.
To address C2, we built a
PRIME+PROBE
attack for
observing the MMU's modication on the state of LLC in
commodity Intel processors. We noticed that the noisy nature
of
PRIME+PROBE
that monitors the entire LLC in each round
of the attack makes it difcult to observe the (faint) MMU
signal, but a more directed and low-noise
EVICT+TIME
attack
that monitors one cache set at a time can reliably detect
the MMU signal. We discuss the details of this attack for
derandomizing JavaScript's heap and code ASLR in Section V.
To address C3, we needed to ensure that we can allocate
and access virtually contiguous buffers that span enough PT
levels to completely derandomize ASLR. For example, on a
64 bit Linux system ASLR entropy for the browser's heap and
the JITed code is 28 bits and on an x8664 processor, there are
4 PT levels, each providing 9 bits of entropy (each PT level
stores 512 PT entries). Hence, we need a virtually contiguous
area that spans all four PT levels for complete derandomization
of ASLR. In Section VI, we discuss how ASLR

Cache
exploits low-level memory management properties of Chrome
and Firefox to gain access to such areas.
IV. T
IMING BY
C
OUNTING
Recent work shows that timing side channels can be
exploited in the browser to leak sensitive information such
as randomized pointers [6] or mouse movement [48]. These
attacks rely on the precise JavaScript timer in order to tell the
difference between an access that is satised through a cache or
main memory. In order to thwart these attacks, major browser
vendors have reduced the precision of the timer. Based on our
measurements, both Firefox and Chrome have decreased the
precision of
performance.now()
to exactly 5

s.
We designed a small microbenchmark in order to bet-
ter understand the quality of the JavaScript timer (i.e.,
Fig. 4. 1. How the old
performance.now()
was used to distinguish be-
tween a cached or a memory access. 2. How the new
performance.now()
stops timing side-channel attacks. 3. How the SMC can be used to make
the distinction in the memory reference using a separate counting core as a
reference. 4. How TTT can make the distinction by counting in between ticks.
performance.now()
) in our target browsers. The mi-
crobenchmark measures how many times we can execute
performance.now()
in a tight loop between two sub-
sequent ticks of
performance.now()
for a hundred
times. Figure 3 shows the results of our experiment in
terms of frequency. Firefox shows a single peak, while
Chrome shows multiple peaks. This means that Firefox's
performance.now()
ticks exactly at 5

s, while Chrome
has introduced some jitter around around the 5

s intervals.
The decreased precision makes it difcult to tell the difference
between a cached or memory access (in the order of tens of
nanoseconds) which we require for AnC to work.
Figure 4.1 shows how the old timer was being used to
distinguish between cached or memory access (CT stands for
cache time and MT stands for memory time). With a low-
precision timer, shown in Figure 4.2, it is no longer possible to
tell the difference by simply calling the timer. Recent work [35]
shows that it is possible to use the clock edge to improve the
precision of the degraded counters, but it is still not enough
to tell the difference between operations that are only tens of
nanoseconds apart (i.e., accessing cache versus memory).
In the following sections, we describe two techniques for
measuring how long executing a memory reference takes by
counting how long a memory reference takes rather than
timing. Both techniques rely on the fact that CPU cores have
a higher precision than
performance.now()
.
The rst technique (Figure 4.3), shared memory counter
(SMC), relies on an experimental feature (with a draft
RFC [55]) that allows for sharing of memory between
JavaScript's web workers
2
. SMC builds a high-resolution
counter that can be used to reliably implement AnC in all
the browsers that implement it. Both Firefox and Chrome
currently support this feature, but it needs to be explicit
enabled due to its experimental nature. We expect shared
memory between JavaScript web workers to become a default-
on mainstream feature in a near future. The second technique2
JavaScript webworkers are threads used for long running background tasks.
5

--- page 8 ---

0 0.1 0.2 0.3 0.4 0.5 0.6 0 20 40 60 80 100 120 140Frequency (normalized)Number of loops per tick of performance.now()
Google Chrome 50.0 on Linux 4.4.0Mozilla Firefox 46.0.1 on Linux 4.4.0

--- page 9 ---

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--- page 10 ---

(Figure 4.4), time to tick (TTT), relies on the current low-
precision
performance.now()
for building a timer that
allows us to measure the difference between a cached reference
and a memory reference, and allows us to implement AnC in
low-jitter browsers such as Firefox.
The impact of TTT and SMC goes beyond AnC.
All previ-
ous timing attacks that were considered mitigated by browser
vendors are still applicable today.
It is worth mentioning that
the recently proposed fuzzy time defense for browsers [35],
while also expensive, is not effective against SMC. We now
discuss the TTT and SMC timers in further detail.
A. Time to Tick
The idea behind the TTT measurement, as shown in
Figure 4.4, is quite simple. Instead of measuring how long
a memory reference takes with the timer (which is no longer
possible), we count how long it takes for the timer to tick
after
the memory reference takes place. More precisely, we
rst wait for
performance.now()
to tick, we then ex-
ecute the memory reference, and then count by executing
performance.now()
in a loop until it ticks. If memory
reference is a fast cache access, we have time to count more
until the next tick in comparison to a memory reference that
needs to be satised through main memory.
TTT performs well in situations where
performance.now()
does not have jitter and ticks
at regular intervals such as in Firefox. We, however, believe
that TTT can also be used in
performance.now()
with
jitter as long as it does not
drift
, but it will require a higher
number of measurements to combat jitter.
B. Shared Memory Counter
Our SMC counter uses a dedicated JavaScript web worker
for counting through a shared memory area between the main
JavaScript thread and the counting web worker. This means
that during the attack, we are practically using a separate core
for counting. Figure 4-3 shows how an attacker can use SMC
to measure how long a memory reference takes. The thread
that wants to perform the measurement (in our case the main
thread) reads the counter and stores it in
c
1
, executes the
memory reference, and then reads the counter again and stores
it in
c
2
. Since the other thread is incrementing the counter
during the execution of the memory reference, in case of a
slow memory access, we see a larger
c
2

c
1
compared to the
case where a faster cache access is taking place.
SMC is agnostic to the quality of
performance.now()
since it only relies on a separate
observer
core for its measure-
ments.
C. Discussion
We designed a microbenchmark that performs a cached
access and a memory access for a given number of iterations.
We can do this by accessing a huge buffer (an improvised
eviction set), ensuring the next access of a test buffer will be
uncached. We measure this access time with both timers. We
then know the next access time of the same test buffer will
be cached. We time this access with both timers. In all cases,
TTT and SMC could tell the difference between the two cases.
We use TTT on Firefox and SMC on Chrome for the
evaluation of AnC in Section VII. The shared memory feature,
needed for SMC, is currently enabled by default in the nightly
build of Firefox, implemented in Chrome [10] and under
development in Edge [8]. We have notied major browser
vendors warning them of this danger.
V. I
MPLEMENTING
A
N
C
Equipped with our TTT and SMC timers, we now proceed
with the implementation of AnC described in Section III-D.
We rst show how we managed to trigger MMU walks when
accessing our target heap and when executing code on our
target JIT area in Section V-A. We then discuss how we
identied the page offsets that store PT entries of a target
virtual address in Sections V-B, V-C and V-D. In Sections V-E
and V-F, we describe the techniques that we applied to observe
the signal and uniquely identify the location of PT entries
inside the cache lines that store them. In Sections V-G and V-H
we discuss the techniques we applied to clear the MMU signal
by ushing the page table caches and eliminating noise.
A. Triggering MMU Page Table Walks
In order to observe the MMU activities on the CPU caches
we need to make sure that 1) we know the offset in pages
within our buffer when we access the target, and 2) we are
able to evict the TLB in order to trigger an MMU walk on the
target memory location. We discuss how we achieved these
goals for heap memory and JITed code.
1) Heap:
We use the
ArrayBuffer
type to back
the heap memory that we are trying to derandomize. An
ArrayBuffer
is always page-aligned which makes it pos-
sible for us to predict the relative page offset of any index in
our target
ArrayBuffer
. Recent Intel processors have two
levels of TLB. The rst level consists of an instruction TLB
(iTLB) and a data TLB (dTLB) while the second level is a
larger unied TLB cache. In order to ush both the data TLB
and the unied TLB, we access every page in a TLB eviction
buffer with a larger size than the unied TLB. We later show
that this TLB eviction buffer can be used to evict LLC cache
sets at the desired offset as well.
2) Code:
In order to allocate a large enough JITed code
area we spray
2
17
JavaScript functions in an
asm.js
[26]
module. We can tune the size of these functions by changing
the number of their statements to be compiled by the JIT
engine. The machine code of these functions start from a
browser-dependent but known offset in a page and follow each
other in memory and since we can predict their (machine code)
size on our target browsers, we know the relative offset of each
function from the beginning of the
asm.js
object. In order
to minimize the effect of these functions on the cache without
affecting their size, we add an
if
statement in the beginning
of all the functions in order not to execute their body. The goal
is to hit a single cache line once executed so as to not obscure
the pagetable cache line signals, but still maintain a large offset
between functions. To trigger a PT walk when executing one
of our functions, we need to ush the iTLB and the unied
TLB. To ush the iTLB, we use a separate
asm.js
object and
execute some of its functions that span enough pages beyond
the size of the iTLB. To ush the unied TLB, we use the
same TLB eviction buffer that we use for the heap.
6

--- page 11 ---

As we will discuss shortly, AnC observes one page offset
in each round. This allows us to choose the iTLB eviction
functions and the page offset for the unied TLB eviction
buffer in a way that does not interfere with the page offset
under measurement.
B. PRIME+PROBE and the MMU Signal
The main idea behind ASLR

Cache is the fact that we
can observe the effect of MMU's PT walk on the LLC.
There are two attacks that we can implement for this pur-
pose [49]:
PRIME+PROBE
or
EVICT+TIME
. To implement a
PRIME+PROBE
attack, we need to follow a number of steps:
1) Build optimal LLC eviction sets for all available page
colors. An optimal eviction set is the precise number
of memory locations (equal to LLC set-associativity)
that once accessed, ensures that a target cache line has
been evicted from the LLC cache set which hosts the
target cache line.
2) Prime the LLC by accessing all the eviction sets.
3) Access the target virtual address that we want to
derandomize, bringing its PT entries into LLC.
4) Probe the LLC by accessing all the eviction sets and
measure which ones take longer to execute.
The eviction sets that take longer to execute presumably
need to fetch one (or more) of their entries from memory. Since
during the prime phase, the entries in the set have been brought
to the LLC, and the only memory reference (besides TLB
eviction set) is the target virtual address, four of these “probed”
eviction sets have hosted the PT entries for the target virtual
address. As we mentioned earlier, these cache sets uniquely
identify the upper six bits of the PT entry offset at each PT
level.
There are, however, two issues with this approach. First,
building optimal LLC eviction sets from JavaScript, necessary
for
PRIME+PROBE
, while has recently been shown to be pos-
sible [48] takes time, specially without a precise timer. Second,
and more fundamental, we cannot perform the
PRIME+PROBE
attack reliably, because the very thing that we are trying
to measure, will introduce noise in the measurements. More
precisely, we need to ush the TLB before accessing our
target virtual address. We can do this either before or after
the priming step, but in either case evicting the TLB will
cause the MMU to perform some unwanted PT walks. Assume
we perform the TLB eviction
before
the prime step. In the
middle of accessing the LLC eviction sets during the prime
step, potentially many TLB misses will occur, resulting in
PT walks that can potentially ll the already primed cache
sets, introducing many false positives in the probe step. Now
assume we perform the TLB eviction step
after
the prime step.
A similar situation happens: some of the pages in the TLB
eviction set will result in a PT walk, resulting in lling the
already primed cache sets and again, introducing many false
positives in the probe step.
Our initial implementation of AnC used
PRIME+PROBE
.
It took a long time to build optimal eviction sets and ultimately
was not able to identify the MMU signal due to the high
ratio of noise. To resolve these issues, we exploited unique
properties of our target in order not to build optimal eviction
sets (Section V-C), and due to the ability to control the
trigger (MMU's PT walk), we could opt for a more exotic
EVICT+TIME
attack that allowed us to avoid the drawbacks
of
PRIME+PROBE
(Section V-D).
C. Cache Colors Do Not Matter for AnC
Cache-based side-channel attacks benet from the ne-
grained information available in the state of cache after a secret
operation—the cache sets that were accessed by a victim. A
cache set is uniquely identied by a color (i.e., page color)
and a page (cache line) offset. For example, a cache set in an
LLC with 8192 cache sets can be identied by a (color, offset)
tuple, where
0

color
<
128
and
0

oset
<
64
.
ASLR encodes the secret (i.e., the randomized pointer) in
the page offsets. We can build one eviction set for each of the
64 cache line offsets within a page, evicting all colors of that
cache line offset with each set. The only problem is that the
PT entries at different PT levels may use different page colors,
and hence, show us overlapping offset signals. But given that
we can control the observed virtual address, relative to our
target virtual address, we can control PT entry offsets within
different PT levels as discussed in Section III-D to resolve this
problem.
Our
EVICT+TIME
attack, which we describe next, does
not rely on the execution time of eviction sets. This means that
we do not require to build optimal eviction sets. Coupled with
the fact that ASLR is agnostic to color, we can use
any
page
as part of our eviction set. There is no way that page tables
might be allocated using a certain color layout scheme to avoid
showing this signal, as all of them appear in our eviction sets.
This means that with a sufciently large number of memory
pages, we can evict any PT entry from LLC (and L1D and
L2) at a given page offset, not relying on optimal eviction sets
that take a long time to build.
D. EVICT+TIME Attack on the MMU
The traditional side-channel attacks on cryptographic keys
or for eavesdropping benet from observing the state of the
entire LLC. That is the reason why side-channel attacks
such as
PRIME+PROBE
[48] and
FLUSH+RELOAD
[64] that
allow attackers to observe the entire state of the LLC are
popular [27], [30], [31], [39], [49], [66].
Compared to these attacks,
EVICT+TIME
can only
gain information about one cache set at each measurement
round, reducing its bandwidth compared to attacks such as
PRIME+PROBE
[49].
EVICT+TIME
further makes a strong
assumption that the attacker can observe the performance of
the victim as it performs the secret computation. While these
properties often make
EVICT+TIME
inferior compared to
more recent cache attacks, it just happens that it easily applies
to AnC: AnC does not require a high bandwidth (e.g., to break
a cryptographic key) and it can monitor the performance of the
victim (i.e., the MMU) as it performs the secret computation
(i.e., walking the PT).
EVICT+TIME
implements AnC in the
following steps:
1) Take a large enough set of memory pages to act as
an eviction set.
2) For a target cache line at offset
t
out of the possible
64 offsets, evict that cache line by reading the same
7

--- page 12 ---

Fig. 5. The MMU memorygram as we access the target pages in a pattern that allows us to distinguish between different PT signals. Each row shows MMU
activity for one particular page within our buffer. The activity shows different cache lines within the page table pages that are accessed during the MMU
translation of this page. The color is brighter with increased latency when there is MMU activity. We use different access patterns within our buffer to distinguish
between signals of PT entries at different PT levels. For example, the stair case (on the left) distinguishes the PT entry at PTL1 since we are accessing pages
that are 32 KB apart in succession (32 KB is 8 PT entries at PTL1 or a cache line). Hence, we expect this access pattern to make the moving PTL1 cache line
create a stair case pattern. Once we identify the stair case, it tells us the PT entry slot in PTL1 and distinguishes PTL1 from the other levels. Once a sufciently
unique solution is available for both code and data accesses at all PT levels, AnC computes the (derandomized) 64 bit addresses for code and data, as shown.
offset in all the memory pages in the eviction set.
Accessing this set also ushes the dTLB and the
unied TLB. In case we are targeting code, ush the
iTLB by executing functions at offset
t
.
3) Time the access to the target virtual address that we
want to derandomize at a different cache line offset
than
t
, by dereferencing it in case of the heap target
or executing the function at that location in case of
the code target.
The third step of
EVICT+TIME
triggers a PT walk and
depending on whether
t
was hosting a PT entry cache line, the
operation will take longer or shorter.
EVICT+TIME
resolves
the issues that we faced with
PRIME+PROBE
: rst, we do not
need to create optimal LLC eviction sets, since we do not rely
on eviction sets for providing information and second, the LLC
eviction set is unied with the TLB eviction set, reducing noise
due to fewer PT walks. More importantly, these PT walks (due
to TLB misses) result in signicantly fewer false positives,
again because we do not rely on probing eviction sets for
timing information.
Due to these improvements, we could observe cache line
offsets corresponding to the PT entries of the target virtual ad-
dress when trying
EVICT+TIME
on all 64 possible cache line
offsets in JavaScript both when dereferencing heap addresses
and executing JIT functions. We provide further evaluation in
Section VII, but before that, we describe how we can uniquely
identify the offset of the PT entries inside the cache lines
identied by
EVICT+TIME
.
E. Sliding PT Entries
At this stage, we have identied the (potentially overlap-
ping) cache line offsets of the PT entries at different PT levels.
There still remains two sources of entropy for ASLR: it is
not possible to distinguish which cache line offset belongs to
which PT level, and the offset of the PT entry within the cache
line is not yet known. We address both sources of entropy
by allocating a sufciently large buffer (in our case a 2 GB
allocation) and accessing different locations within this buffer
in order to derandomize the virtual address where the buffer
has been allocated from. We derandomize PTL1 and PTL2
differently than how we derandomize PTL3 and PTL4. We
describe both techniques below.
1) Derandomizing PTL1 and PTL2:
Let's start with the
cache line that hosts the PT entry at PTL1 for a target virtual
address
v
. We observe when one of the (possible) 4 cache line
change as we access
v
+
i

4
KB for
i
=
f
1
;
2
; : : : ;
8
g
. If one
of the cache lines changes at
i
, it immediately provides us with
two pieces of information: the changed cache line is hosting
the PT entry for PTL1 and the PTL1's PT entry offset for
v
is
8

i
. We can perform the same technique for derandomizing
the PT entry at PTL2, but instead of increasing the address by
4 KB each time, we now need to increase by 2 MB to observe
the same effect for PTL2. As an example, Figure 5 shows an
example MMU activity that AnC observes as we change the
cache line for the PT entry at PTL1.
2) Derandomizing PTL3 and PTL4:
As we discussed in
Section III-E, in order to derandomize PTL3, we require an
8 GB crossing in the virtual address space within our 2 GB
allocation and to derandomize PTL4, we require a 4 TB virtual
address space crossing to happen within our allocation. We rely
on the behavior of memory allocators, discussed in Section VI,
in the browsers to ensure that one of our (many) allocations
satises this property. But assuming that we have a cache
line change at PTL3 or PTL4, we would like to detect and
derandomize the corresponding level. Note that a cache line
crossing at PTL4 will inevitably cause a cache line crossing
at PTL3 too.
Remember that each PT entry at PTL3 covers 1 GB of
virtual memory. Hence, if a cache line crossing at PTL3
happens within our 2 GB allocation, then our allocation could
cover either two PTL3 PT entries, when crossing is exactly at
the middle of our buffer, or three PT entries. Since a crossing
8

--- page 13 ---

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--- page 14 ---

exactly in the middle is unlikely, we consider the case with
three PT entries. Either two of the three or one of three PT
entries are in the new cache line. By observing the PTL3 cache
line when accessing the rst page, the middle page, and the
last page in our allocation, we can easily distinguish between
these two cases and fully derandomize PTL3.
A cache line crossing at PTL4 only occurs if the cache
line at PTL3 is in the last slot in its respective PT page. By
performing a similar technique (i.e., accessing the rst and last
page in our allocation), if we observe a PT entry cache line
PTE
2
change from the last slot to the rst slot and another PT
entry cache line PTE
1
move one slot ahead, we can conclude
a PTL4 crossing and uniquely identify PTE
2
as the PT entry
at PTL3 and PTE
1
as the PT entry at PTL4.
F. ASLR Solver
We created a simple solver in order to rank possible
solutions against each other as we explore different pages
within our 2 GB allocation. Our solver assumes 512 possible
PT entries for each PT level for the rst page of our allocation
buffer, and ranks the solutions at each PT level independently
of the other levels.
As we explore more pages in our buffer according to
patterns that we described in Section V-E1 and Section V-E2,
our solver gains a signicant condence in one of the solutions
or gives up and starts with a new 2 GB allocation. A solution
will always derandomizes PTL1 and PTL2 and also PTL3 and
PTL4 if there was a cache line crossing at these PT levels.
G. Evicting Page Table Caches
As mentioned in Section III-B, some processors may cache
the translation results for different page table levels in their
TLB. AnC needs to evict these caches in order to observe the
MMU signal from all PT levels. This is straightforward: we
can access a buffer that is larger than that the size of these
caches as part of our TLB and LLC eviction.
For example, a Skylake i7-6700K core can cache 32 entries
for PTL2 look ups. Assuming we are measuring whether there
is page table activity in the
i
-th cache line of page table pages,
accessing a 64 MB (i.e., 32

2 MB) buffer at 0 +
i

64
, 2 MB
+
i

64
, 4 MB +
i

64
, . . . , 62 MB +
i

64
will evict the
PTL2 page table cache.
While we needed to implement this mechanism natively
to observe the signal on all PT levels, we noticed that due
to the JavaScript runtime activity, these page table caches are
naturally evicted during our measurements.
H. Dealing with Noise
The main issue when implementing side-channel attacks is
noise. There are a number of countermeasures that we deploy
in order to reduce noise. We briey describe them here:
1) Random exploration:
In order to avoid false negatives
caused by the hardware prefetcher, we select
t
(the page offset
that we are evicting) randomly within the possible remaining
offsets that we (still) need to explore. This randomization also
helps by distributing the localized noise caused by system
events.
2) Multiple rounds for each offset:
In order to add reli-
ability to the measurements, we sample each offset multiple
times ('rounds') and consider the median for deciding a cached
versus memory access. This simple strategy reduces the false
positives and false negatives by a large margin. For large scale
experiment and visualization on the impact of measurement
rounds vs other solving parameters, please see section VII-C.
For an AnC attack, false negatives are harmless due to the
fact that the attacker can always retry with a new allocation as
we discuss in the next section. We evaluate the success rate,
false positives and false negatives of the AnC attack using
Chrome and Firefox in Section VII.
I. Discussion
We implemented two versions of AnC. A native implemen-
tation in C in order to study the behavior of the MMU PT walk
activity without the JavaScript interference and a JavaScript-
only implementation.
We ported the native version to different architectures and
Microsoft Windows 10 to show the generality of AnC pre-
sented in Section VII-D. Our porting efforts revolved around
implementing a native version of SMC (Section IV-B) to
accurately differentiate between cached and uncached memory
accesses on ARM which only provides coarse-grained (0.5

s)
timing mechanism and dealing with different page table struc-
tures. Our native implementation amounts to 1283 lines of
code
Our JavaScript-only implementation works on Chrome and
Firefox browsers and is aimed to show the real-world impact
of the AnC attack presented in various experiments in Sec-
tion VII. We needed to handtune the JavaScript implementation
using
asm.js
[26] in order to make the measurements faster
and more predictable. This limited our allocation size to the
maximum of 2 GB. Our JavaScript implementation amounts to
2370 lines of code.
VI. A
LLOCATORS AND
A
N
C
As mentioned in Section V-E2, we rely on the behavior of
the memory allocators in the browsers to get an allocation that
crosses PTL3 and PTL4 in the virtual address space. We briey
discuss the behavior of the memory allocators in Firefox and
Chrome and how we could take advantage of them for AnC.
A. Memory Allocation in Firefox
In Firefox, memory allocation is based on demand paging.
Large object allocations from a JavaScript application in the
browser's heap is backed by
mmap
without
MAPPOPULATE
.
This means that memory is only allocated when the corre-
sponding page in memory is touched.
Figure 6 shows how Firefox's address space is laid out in
memory. Firefox uses the stock
mmap
provided by the Linux
kernel in order to randomize the location of JITed code and
heap using 28 bits of entropy. The (randomized) base address
for
mmap
is only chosen once (by the OS) and after that the
subsequent allocations by Firefox grow backward from the
previous allocation towards low (virtual) memory. If an object
is deleted, Firefox reuses its virtual memory for the subsequent
allocations. Hence, to keep moving backward in the virtual
9

--- page 15 ---

Fig. 6. Low-level memory allocation strategies in Firefox and Chrome.
Firefox uses the stock
mmap
in order to gain ASLR entropy for its JITed code
and heap while Chrome does not rely on
mmap
for entropy and randomizes
each large code/data allocation.
address space, a JavaScript application should linger to its old
allocations.
An interesting observation that we made is that a JavaScript
application can allocate TBs of (virtual) memory for its objects
as long as they are not touched. AnC exploits this fact and
allocates a few 2 GB buffers for forcing a cache line change at
PTL3 (i.e., 1 bit of entropy remaining), or if requested, a large
number of 2 GB objects forcing a cache line change at PTL4
(i.e., fully derandomized).
To obtain a JIT code pointer, we rely on our heap pointer
obtained in the previous step. Firefox reserves some virtual
memory in between JIT and heap. We rst spray a number
of JITed objects to exhaust this area right before allocating
our heap. This ensures that our last JITed object is allocated
before our heap. There are however a number of other objects
that JavaScript engine of Firefox allocates in between our last
JITed object and the heap, introducing additional entropy. As
a result, we can predict the PTL3 and PTL4 slots of our target
JIT pointer using our heap pointer, but the PTL1 and PTL2
slots remain unknown. We now deploy our code version of
AnC to nd PTL1 and PTL2 slots of our code pointer, resulting
in a full derandomization.
B. Memory Allocation in Chrome
In Chrome, memory allocations are backed by
mmap
and
initialized. This means that every allocation of a certain size
will consume the same amount of physical memory (plus
a few pages to back its PT pages). This prohibits us from
using multiple allocations similar to Firefox. Chrome internally
chooses the randomized location for
mmap
and this means that
for every new large object (i.e., a new heap). This allows for
roughly 35 bits out of the available 36 bits of entropy provided
by hardware (Linux kernel is always mapped in the upper part
of the address space). Randomizing every new heap is devised
in order to protect against the exploitation of use-after-free
bugs that often rely on predictable reuse on the heap [57].
AnC exploits this very protection in order to acquire an
object that crosses a PTL3 or a PTL4 cache line. We rst
Fig. 7. The success rate, false positive and false negative rate of AnC.
allocate a buffer and use AnC to see whether there are PTL3
or PTL4 cache line crossings. If this is not the case, we delete
the old buffer and start with a new allocation. Based on a
given probability
p
, AnC's
i
th allocation will cross PTL3 based
on the following formula using a Bernoulli trial (assuming a
2 GB allocation):
P
i
1
14
(
34
)
i

p
. Calculating for average (i.e.,
p
= 0
:
5
), AnC requires around 6.5 allocations to get a PTL3
crossing. Solving the same equation for a PTL4 crossing, AnC
requires on average 1421.2 allocations to get a crossing. In
a synthetic experiment with Chrome, we observed a desired
allocation after 1235 trials. While nothing stops AnC from
derandomizing PTL4, the large number of trials makes it less
attractive for attackers.
This technique works the same for allocations of both heap
and JITed objects. The current version of AnC implements
derandomization of heap pointers on Chrome using this tech-
nique.
VII. E
VALUATION
We show the success rate and feasibility of the AnC attack
using Firefox and Chrome. More concretely, we like to know
the success rate of AnC in face of noise and the speed in which
AnC can reduce the ASLR entropy. We further compare AnC
with other software-based attacks in terms of requirements and
performance and showcase an end-to-end exploit using a real
vulnerability with pointers leaked by AnC. For the evaluation
of the JavaScript-based attacks, we used an Intel Skylake i7-
6700K processor with 16 GB of main memory running Ubuntu
16.04.1 LTS as our evaluation testbed. We further show the
generality of the AnC attack using various CPU architectures
and operating systems.
A. Success Rate
To evaluate the reliability of AnC, we ran 10 trials of
the attack on Firefox and Chrome and report success rate,
false positive and false negative for each browser. For the
ground truth, we collected run time statistics from the virtual
mappings of the browser's process and checked whether our
guessed addresses indeed match them. In case of a match,
the trial counts towards the success rate. In case AnC fails
to detect a crossing, that trial counts towards false negatives.
False negatives are not problematic for AnC, since it results
in a retry which ultimately makes the attack take longer.
False positives, however, are problematic and we count them
when AnC reports an incorrect guessed address. In case of
Chrome, we report numbers for when there are PTL3 cache
line crossings. We did not observe a PTL4 crossing (i.e., all
levels) in our trials. In case of Firefox, we performed the
measurement by restarting it to get a new allocation each time
10

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--- page 17 ---

0 0.2 0.4 0.6 0.8 1Chrome 3 LevelsFirefox 3 LevelsFirefox 4 LevelsFalse positiveFalse negativeSuccess rate

--- page 18 ---

Fig. 8. Reduction in heap (Chrome and Firefox) and JIT (Firefox) ASLR
entropy over time with the AnC attack. At the code stage of the attack, AnC
already knows the exact PTE slots due to the obtained heap pointer. This
means that for the code, the entropy reaches zero at 37.9 s, but our ASLR
solver is agnostic to this information.
and we used it to derandomize both JIT and heap pointers. We
report numbers for both PTL3 and PTL4 cache line crossings.
Figure 7 reports the results of the experiment. In the case
of Chrome, AnC manages to successfully recover 33 bits out
of the 35 bits of the randomized heap addresses in 8 of the
cases. In the case of Firefox, AnC reduces the entropy of JIT
and heap to a single bit in all 10 cases. Getting the last bit
is successful in 6 of the cases with 2 cases as false positive.
The PTE hosting the last bit of entropy (i.e., in PTL4) is often
shared with other objects in the Firefox runtime, making the
measurements more noisy compared to PTEs in other levels.
B. Feasibility
To evaluate the feasibility of AnC from an attacker's
perspective, we report on the amount of time it took AnC
to reduce ASLR's entropy in the same experiment that we just
discussed. Allocating the buffers that we use for the AnC do
not take any time on Chrome. On Firefox, for crossing a cache
line in PTL3, the buffer allocations take 5.3 s, and for crossing
a cache line in PTL4, the buffer allocations take 72.7 s.
Figure 8 shows ASLR entropy as a function of time when
AnC is applied to Firefox and Chrome as reported by our
solver described in Section V-F. Both heap and code deran-
domization are visualized. Note that due to the noise our solver
sometimes needs to consider more solutions as time progress
resulting in a temporary increase in the estimated entropy.
More importantly, our solver is agnostic to the limitations of
the underlying ASLR implementations and always assumes
36 bits of entropy (the hardware limit). This means that AnC
can reduce the entropy even if the implementation uses all
available bits for entropy which is not possible in practice. In
the case of Chrome, in 11.2 s the entropy is reduced to only
2 bits (our solver does not know about kernel/user space split
and reports 3 bits). In the case of Firefox, in 33.1 s the entropy
is reduced to 1 bit when crossing a cache line in PTL3 (our
solver does not know about
mmap
's entropy) and in 40.5 s to
zero when crossing a cache line in PTL4.
As discussed in Section VI, after obtaining a heap pointer,
our AnC attack proceeds to obtain a JITed code pointer. At
Fig. 9. The effects of our noise reduction techniques on the delity of
AnC. The plotted intensity indicates false positive (wrong answer) rate, as a
function of solver condence requirement (X axis) vs. number of measurement
repetitions (Y axis). This shows that the number of measurement rounds
is critical to reliable conclusions while the condence margin improves the
results further.
this stage of the attack, AnC already knows the upper two
PT slots of the JIT area (our solver does not know about
this information). After 37.9 s, AnC reduces the code pointer
entropy to only 6 bits as reported by our ASLR solver. These
are the same entropy bits that are shared with our heap pointer.
This means that at this stage of the attack we have completely
derandomized code and heap ASLR in Firefox.
C. Noise
We evaluated the efcacy of our techniques against noise
in the system. As we mentioned earlier, we used measurement
rounds in order to combat temporal noise and a scoring system
in our solver in order to combat more persistent noise in the
system.
Figure 9 shows different conguration of AnC with respect
to the number of rounds and the condence margin in our
solver. As we decrease the number of rounds or condence
margin, we observe more false positives in the system. These
results show that with our chosen conguration (condence
margin = 10 and rounds = 20) these techniques are effective
against noise.
D. Generalization
Our evaluation so far shows that AnC generalizes to
different browsers. We further studied the generality of the
AnC attack by running our native implementation on different
CPU architectures.
Table I shows a successful AnC attack on 11 different
CPU architectures including Intel, ARM and AMD. We did
not nd an architecture on which the AnC attack was not
possible. Except for on ARMv7, we could fully derandomize
ASLR on all architectures. On ARMv7, the top level page
table spans four pages which introduces two bits of entropy in
the high virtual address bits. On ARMv7 with physical address
extension (i.e., ARMv7+LPAE), there are only four entries in
the top level page table which t into a cache line, resulting
again in two remaining bits of entropy. On ARMv8, AnC fully
solves ASLR given a similar page table structure to x8664.
11

--- page 19 ---

0 3 6 9 12 15 18 21 24 27 30 33 36 0 10 20 30 40 50Remaining virtual address entropy (bits)Elapsed time (s)
Chrome heap (PTL3 cacheline crossing)Firefox heap (PTL3 cacheline crossing)Firefox heap (PTL4 cacheline crossing)Firefox JIT (PTL2 cacheline crossing)

--- page 20 ---

0 4 8 12 16 20 24 28 32 36 40 0 2 4 6 8 10 12 14 16 18 20 22Measurement roundsCondence margin
False positives out of 10
False positives out of 10 0 2 4 6 8 10

--- page 21 ---

TABLE I. CPU
MODELS VERIFIED TO BE AFFECTED BY
A
N
C.
Vendor CPU Model Year MicroarchitectureIntel Core i7-6700K 2015 Skylake
Intel Core i3-5010U 2015 Broadwell
Allwinner A64 2015 ARMv8-A, Cortex-A53
Nvidia Jetson TK-1 2014 ARMv7, Cortex-A15
Nvidia Tegra K1 CD570M 2014 ARMv7+LPAE, Cortex-A15
Intel Core i7-4510U 2014 Haswell
Intel Celeron N2840 2014 Silvermont
Intel Atom C2750 2013 Silvermont
AMD FX-8320 8-Core 2012 Piledriver
Intel Core i7-3632QM 2012 Ivy Bridge
Intel E56xx/L56xx/X56xx 2010 Westmere
Both ARM and AMD processors have exclusive LLCs
compared to Intel. These results show that AnC is agnostic to
the inclusivity of the LCC. We also successfully performed the
AnC attack on the Microsoft Windows 10 operating system.
E. Comparison against Other Derandomization Attacks
Table II compares existing derandomization attacks against
ASLR with AnC. Note that all the previous attacks rely on
software features that can be mitigated. For example, the
most competitive solution, Dedup Est Machina [6], relies on
memory deduplication, which was only available natively on
Windows and has recently been turned off by Microsoft [14],
[46]. Other attacks require crash-tolerance or crash-resistance
primitives, which are not always available and also yield much
more visible side effects. We also note that AnC is much faster
than all the existing attacks, completing in 150 seconds rather
than tens of minutes.
F. End-to-end attacks
Modern browsers deploy several defenses such as ASLR or
segment heaps to raise the bar against attacks [63]. Thanks to
such defenses, traditional browser exploitation techniques such
as heap spraying are now much more challenging to execute,
typically forcing the attacker to derandomize the address space
before exploiting a given vulnerability [53].
For example, in a typical vtable hijacking exploit (with
many examples in recent Pwn2Own competitions), the attacker
seeks to overwrite a vtable pointer to point to a fake vtable
using type confusion [38] or other software [60] or hard-
ware [6] vulnerabilities. For this purpose, the attacker needs
to leak code pointers to craft the fake vtable and a heap
pointer to the fake vtable itself. In this scenario, nding a
dedicated information disclosure primitive is normally a sine
qua non to mount the attack. With AnC, however, this is no
longer a requirement: the attacker can directly leak heap and
code addresses she controls with a cache attack against the
MMU. This signicantly reduces the requirements for end-to-
end attacks in the info leak era of software exploitation [53].
As an example, consider CVE-2013-0753, a use-after-free
vulnerability in Firefox. An attacker is able to overwrite a
pointer to an object and this object is later used to perform
a virtual function call, using the rst eld of the object to
reference the object's vtable. On 32-bit Firefox, this vulner-
ability can be exploited using heap spraying, as done in the
publicly available Metasploit module (https://goo.gl/zBjrXW).
However, due to the much larger size of the address space,
an information disclosure vulnerability is normally required
TABLE II. D
IFFERENT ATTACKS AGAINST USER
-
SPACE
ASLR.
Attack Time Probes Pointers RequirementBROP [5] 20 m
4000
Code Crash tolerance
CROP [19] 243 m
2
28
Heap/code Crash resistance
Dedup Est Machina [6] 30 m
0
Heap/code Deduplication
AnC 150 s
0
Heap/code Cache
on 64-bit Firefox. With AnC, however, we re-injected the
vulnerability in Firefox and veried an attacker can mount an
end-to-end attack without an additional information disclosure
vulnerability. In particular, we were able to (i) craft a fake
vtable containing AnC-leaked code pointers, (ii) craft a fake
object pointing to the AnC-leaked address of the fake vtable,
(iii) trigger the vulnerability to overwrite the original object
pointer with the AnC-leaked fake object pointer, and (iv) hijack
the control ow.
G. Discussion
We showed how AnC can quickly derandomize ASLR
towards end-to-end attacks in two major browsers with high
success rate. For example, AnC can derandomize 64 bit code
and heap pointers
completely
in 150 seconds in Firefox. We
also showed that AnC has a high success rate.
VIII. I
MPACT ON
A
DVANCED
D
EFENSES
The AnC attack casts doubt on some of the recently
proposed advanced defenses in academia. Most notably, AnC
can fully break or signicantly weaken defenses that are
based on information hiding in the address-space and leakage-
resilient code randomization. We discuss these two cases next.
A. Information Hiding
Hiding security-sensitive information (such as code point-
ers) in a large 64 bit virtual address-space in a common
technique to bootstrap more advanced security defenses [9],
[15], [36], [42], [52]. Once the location of the so-called
safe-
region
is known to the attacker, she can compromise the
defense mechanism in order to engage in, for example, control-
ow hijacking attacks [17], [19], [23], [47].
With the AnC attack, in situations where the attacker can
operate arbitrary memory accesses, the unknown (randomized)
target virtual address can be derandomized. Already triggering
a single memory reference in the safe-region allows AnC to
reduce the entropy of ASLR on Linux to
log
2
(2
10

4!) =
10
:
1
bits (1 bit on the PTL4 and 3 bits on other levels) and on
Windows to
log
2
(2
9

4!) = 9
:
4
bits (9 bits on PTL3, PTL2
and PTL1). Referencing more virtual addresses in different
memory pages allows AnC to reduce the entropy further.
For example, the original linear table implementation for
the safe-region in CPI [36], spans
2
42
of virtual address-space
and moves the protected code pointers in this area. This means
that, since the relative address of secret pointers with respect to
each other is known, an attacker can also implement sliding to
nd the precise location of the safe-region using AnC, breaking
CPI. Similarly, the more advanced two-level lookup table or
hashtable versions of CPI [37] for hiding the safe-region will
be prone to AnC attacks by creating a sliding mechanism on
the target (protected) pointers.
12

--- page 22 ---

We hence believe that randomization-based information
hiding on modern cache-based CPU architectures is inherently
prone to cache attacks when the attacker controls memory
accesses (e.g., web browsers). We thereby caution future
defenses not to rely on ASLR as a pivotal building block,
even when problematic software features such as memory
deduplication [6] or memory overcommit [47] are disabled.
B. Leakage-resilient Code Randomization
Leakage-resilient code randomization schemes based on
techniques like XnR and code pointer hiding [1], [7], [13] aim
to provide protection by making code regions execute-only and
the location of target code in memory fully unpredictable. This
makes it difcult to perform code-reuse attacks given that the
attacker cannot directly or indirectly disclose the code layout.
AnC weakens all these schemes because it can nd the
precise memory location of executed code without reading it
(Section VII-B). Like Information Hiding, the execution of a
single function already leaves enough traces in the cache from
the MMU activity to reduce its address entropy signicantly.
IX. M
ITIGATIONS
Detection
It is possible to detect an on-going AnC attack
using performance counters [50]. These types of anomaly-
based defenses are, however, prone to false positives and false
negatives by nature.
Cache coloring
Partitioning the shared LLC can be used to
isolate an application (e.g., the browser) from the rest of the
system [33], [39], but on top of complications in the kernel's
frame allocator [40], it has performance implications both for
the operating system and the applications.
Secure timers
Reducing the accuracy of the timers [35], [43],
[58] makes it harder for attackers to tell the difference between
cached and memory accesses, but this option is often costly
to implement. Further, there are many other possible sources
to craft a new timer. Prior work [11] shows it is hard if not
impossible to remove all of them even in the context of simple
microkernels. This is even more complicated with browsers,
which are much more complex and bloated with features.
Isolated caches
Caching PT entries in a separate cache rather
than the data caches can mitigate AnC. Having a separate cache
just for page table pages is quite expensive in hardware and
adopting such solution as a countermeasure defeats the purpose
of ASLR—providing a low-cost rst line of defense.
The AnC attack exploits fundamental properties of cache-
based architectures, which improve performance by keeping
hot objects in a faster but smaller cache. Even if CPU man-
ufacturers were willing to implement a completely isolated
cache for PT entries, there are other caches in software that
can be exploited to mount attacks similar to AnC. For example,
the operating system often allocates and caches page table
pages on demand as necessary [24]. This optimization may
yield a timing side channel on memory management operations
suitable for AnC-style attacks. Summarizing, we believe that
the use of ASLR is fundamentally insecure on cache-based
architectures and, while countermeasures do exist, they can
only limit, but not eliminate the underlying problem.
X. R
ELATED
W
ORK
A. Derandomizing ASLR in Software
Bruteforcing, if allowed in software, is a well-known tech-
nique for derandomizing ASLR. For the rst time, Shacham
et al. [54] showed that it is possible to systematically deran-
domize ASLR on 32 bit systems. BROP [5] bruteforces values
in the stack byte-by-byte in order to nd a valid 64 bit return
address using a few hundreds of trials. Bruteforcing, however,
is not always possible as it heavily relies on application-
specic behavior [5], [19]. Further, signicant number of
crashes can be used by anomaly detection systems to block
an ongoing attack [54].
A key weakness of ASLR is the fact that large (virtual)
memory allocations reduce its entropy. After a large memory
allocation, the available virtual address-space is smaller for the
next allocation, reducing entropy. This weakness has recently
been exploited to show insecurities in software hardening
techniques that rely on ASLR [23], [47].
Memory deduplication is an OS or virtual machine monitor
feature that merges pages across processes or virtual machines
in order to reduce the physical memory footprint. Writing to
merged pages results in a copy-on-write that is noticeably
slower than writing to a normal. This timing channel has been
recently used to bruteforce ASLR entropy in clouds [2] or
inside browsers [6].
All these attacks against ASLR rely on a aw in software
that allows an attacker to reduce the entropy of ASLR. While
software can be xed to address this issue, the microarchi-
tectural nature of our attack makes it difcult to mitigate.
We hence recommend decommissioning ASLR as a defense
mechanism or a building block for other defense mechanisms.
B. Timing Attacks on CPU Caches
Closest to ASLR

Cache, in terms of timing attacks on
CPU caches is the work by Hund et al. [27] in breaking Win-
dows kernel-space ASLR from a local user-space application.
Their work, however, assumes randomized physical addresses
(instead of virtual addresses) with a few bits of entropy, and
that the attacker has the ability to reference arbitrary virtual
addresses. Similar attacks geared towards breaking kernel-
level ASLR from a controlled process have been recently
documented using the
prefetch
instruction [25], hardware
transactional memory [32] and branch prediction [18]. In our
work, we derandomized high-entropy virtual addresses in the
browser process inside a sandboxed JavaScript. To the best
of our knowledge, AnC is the rst attack that breaks user-
space ASLR from JavaScript using a cache attack, signicantly
increasing the impact of these types of attacks.
Timing side-channel attack on CPU caches have also
been used to leak private information, such as cryptographic
keys, mouse movements, and etc. The
FLUSH+RELOAD
at-
tack [31], [64], [66] leaks data from a sensitive process by
exploiting the timing differences when accessing cached data.
FLUSH+RELOAD
assumes the attacker has access to victims'
code pages either via the shared page cache or some form of
memory deduplication. The
PRIME+PROBE
attack [39], [49]
lifts this requirement by only relying on cache misses from
the attacker's process to infer the behavior of the victim's
13

--- page 23 ---

process when processing secret data. This relaxation makes
it possible to implement
PRIME+PROBE
in JavaScript in the
browser [48], signicantly increasing the impact of cache side-
channel attacks for the Internet users.
While
FLUSH+RELOAD
and
PRIME+PROBE
observe
the change in the state of the entire cache, the older
EVICT+TIME
[49] attack, used for recovering AES keys,
observes the state of one cache set at a time. In situations
where information on the entire state of the cache is necessary,
EVICT+TIME
does not perform as effectively as the other
attacks, but it has a much higher signal to noise ratio since
it is only observing one cache set at a time. We used the
reliability of
EVICT+TIME
for observing the MMU signal.
C. Defending Against Timing Attacks
As we described in Section IX mitigating AnC is difcult.
However, we discuss some attempts for reducing the capability
of the attackers to perform timing side-channel attacks.
At the hardware-level, TimeWarp [43] reduces the delity
of timers and performance counters to make it difcult for
attackers to distinguish between different microarchitectural
events. Stefan et al. [58] implement a new CPU instruction
scheduling algorithm that is indifferent to timing differences
from underlying hardware components, such as the cache, and
is, hence, secure against cache-based timing attacks.
At the software-level, major browsers have reduced the
accuracy of their timers in order to thwart cache attacks
from JavaScript. Concurrent to our efforts, Kohlbrenner and
Shacham [35] show that it is possible to improve the degraded
timers by looking at when the degraded clock ticks and
proposed introducing noise in the timer and the event loop
of JavaScript. In this paper, we show that it is possible to
build more accurate timers than that of Kohlbrenner and
Shacham, necessary for the correct operation of AnC. Our
SMC timer is independent of the JavaScript-provided timer
and the JavaScript runtime, and will thus not be stopped by
the proposed defenses.
Page coloring, in order to partition the shared cache,
is another common technique for defending against cache
side-channel attacks [39]. Kim et al. [33] propose a low-
overhead cache isolation technique to avoid cross-talk over
shared caches. Such techniques could be retrotted to protect
the MMU from side-channel attacks from JavaScript, but as
mentioned in Section IX, they suffer from deployability and
performance problems. As a result, they have not been adopted
to protect against cache attacks in practical settings. By dynam-
ically switching between diversied versions of a program,
Crane et al. [12] change the mapping of program locations
to cache sets, making it difcult to perform cache attacks
on program's locations. Our AnC attack, however, targets the
MMU operations and can already reduce the ASLR entropy
signicantly as soon as one program location is accessed.
D. Other Hardware-based Attacks
Fault attacks are pervasive for extracting secrets from se-
cure processors [21]. Power, thermal and electromagnetic eld
analysis have been used for building covert channels and ex-
tracting cryptographic keys [3], [20], [41]. Recent Rowhammer
attacks show the possibility of compromising the browser [6],
cloud virtual machines [51] and mobile devices [61] using
wide-spread DRAM disturbance errors [34].
XI. C
ONCLUSIONS
In this paper, we described how ASLR is fundamentally
insecure on modern architectures. Our attack relies on the
interplay between the MMU and the caches during virtual
to physical address translation—core hardware behavior that
is central to efcient code execution on modern CPUs. The
underlying problem is that the complex nature of modern
microarchitectures allows attackers with knowledge of the
architecture to craft a carefully chosen series of memory
accesses which manifest timing differences that disclose what
memory is accessed where and to infer all the bits that make
up the address. Unfortunately, these timing differences are
fundamental and reect the way caches optimize accesses in
the memory hierarchy. The conclusion is that such caching
behavior and strong address space randomization are mutually
exclusive. Because of the importance of the caching hierarchy
for the overall system performance, all xes are likely to
be too costly to be practical. Moreover, even if mitigations
are possible in hardware, such as separate cache for page
tables, the problems may well resurface in software. We hence
recommend ASLR to no longer be trusted as a rst line of
defense against memory error attacks and for future defenses
not to rely on it as a pivotal building block.
D
ISCLOSURE
We have cooperated with the National Cyber Security
Centre in the Netherlands to coordinate the disclosure of AnC
to the affected hardware and software vendors. Most of them
acknowledged our ndings and we are closely working with
some of them to address some of the issues raised by AnC.
A
CKNOWLEDGEMENTS
We would like to thank the anonymous reviewers for their
comments. Stephan van Schaik helped us observe the MMU
signal on ARM and AMD processors. This work was supported
by the European Commission through project H2020 ICT-32-
2014 SHARCS under Grant Agreement No. 644571 and by
the Netherlands Organisation for Scientic Research through
grant NWO 639.023.309 VICI Dowsing.
R
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