--- type: Article title: "Loophole: Timing Attacks on Shared Event Loops in Chrome" resource: "https://www.usenix.org/conference/usenixsecurity17/technical-sessions/presentation/vila" tags: [article, webseclist-reference, en, usenix-org] generated: by: webseclist-refs/1 at: "2026-08-19T16:43:39+00:00" status: stable stale_after: 2027-08-19 sources: - id: original resource: "https://www.usenix.org/conference/usenixsecurity17/technical-sessions/presentation/vila" title: "Loophole: Timing Attacks on Shared Event Loops in Chrome" author: Pepe Vila, Boris Köpf - id: capture resource: "https://web.archive.org/web/20170817194503/https://www.usenix.org/conference/usenixsecurity17/technical-sessions/presentation/vila" also_at: - "https://www.usenix.org/system/files/conference/usenixsecurity17/sec17-vila.pdf" - "https://www.usenix.org/sites/default/files/conference/protected-files/usenixsecurity17_slides_vila.pdf" authors: - Pepe Vila - Boris Köpf canonical_url: "" cited_by: - "2016-17.md:90" commit: "" content_sha256: 3d11ab82591d1c1ce02fb1df2f950fd110a58ca9c0d3c4d26437811617becaea depth: full depth_reason: default kind: article language: en licence: unknown original_url: "https://www.usenix.org/conference/usenixsecurity17/technical-sessions/presentation/vila" published: "" publisher: usenix.org publisher_english: "" raw_sha256: 0ecfcfaf0cb797ae408e4ba1e025b92d95fdab1c6758b08b47ac19cd5de40865 retrieved_from: "https://www.usenix.org/system/files/conference/usenixsecurity17/sec17-vila.pdf" retrieved_kind: live retrieved_utc: "2026-08-19T16:43:39+00:00" slug: usenix-org-loophole-timing-attacks-shared-event-loops-chrome snapshot: 20170817194503 title_english: "" translation_file: "" translation_of: "" --- # Loophole: Timing Attacks on Shared Event Loops in Chrome **Loophole: Timing Attacks on Shared Event Loops in Chrome** - Pepe Vila, Boris Köpf, usenix.org. - Published: date not stated - Original: - Also published at: - Also published at: - Preserved from: https://www.usenix.org/system/files/conference/usenixsecurity17/sec17-vila.pdf (live) on 2026-08-19 - Capture timestamp: 20170817194503 - Licence: unknown Rights remain with the original author and publisher. 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Loophole: Timing Attacks on Shared Event Loops in Chrome Pepe Vila, IMDEA Software Institute & Technical University of Madrid (UPM); Boris Köpf, IMDEA Software Institute https://www.usenix.org/conference/usenixsecurity17/technical-sessions/presentation/vila This paper is included in the Proceedings of the 26th USENIX Security Symposium August 16–18, 2017 • Vancouver, BC, Canada ISBN 978-1-931971-40-9 Open access to the Proceedings of the 26th USENIX Security Symposium is sponsored by USENIX Loophole: Timing Attacks on Shared Event Loops in Chrome Pepe Vila*,† and Boris Köpf* * IMDEA Software Institute † Technical University of Madrid (UPM) {pepe.vila, boris.koepf}@imdea.org Abstract the queue and are sequentially dispatched by the con- trol process according to a FIFO policy. A key fea- Event-driven programming (EDP) is the prevalent ture of EDP is that high-latency (or blocking) opera- paradigm for graphical user interfaces, web clients, and tions, such as database or network requests, can be han- it is rapidly gaining importance for server-side and net- dled asynchronously: They appear in the queue only as work programming. Central components of EDP are events signaling start and completion, whereas the block- event loops, which act as FIFO queues that are used by ing operation itself is handled elsewhere. In this way processes to store and dispatch messages received from EDP achieves the responsiveness and fine-grained con- other processes. currency required for modern user interfaces and net- In this paper we demonstrate that shared event loops work servers, without burdening programmers with ex- are vulnerable to side-channel attacks, where a spy pro- plicit concurrency control. cess monitors the loop usage pattern of other processes by enqueueing events and measuring the time it takes for them to be dispatched. Specifically, we exhibit attacks against the two central event loops in Google’s Chrome web browser: that of the I/O thread of the host process, which multiplexes all network events and user actions, and that of the main thread of the renderer processes, Figure 1: Shared event loop. A enqueues multiple short which handles rendering and Javascript tasks. tasks and records the time at which each of them is pro- For each of these loops, we show how the usage pat- cessed. The time difference between two consecutive tern can be monitored with high resolution and low over- tasks reveals whether V has posted tasks in-between, and head, and how this can be abused for malicious purposes, how long they took to execute. such as web page identification, user behavior detection, and covert communication. In this paper we show that EDP-based systems are sus- ceptible to side-channel attacks. The key observation is 1 Introduction that event loops form a resource that can be shared be- tween mutually distrusting programs. Hence, contention Event-driven programming (EDP) consists of defining of this resource by one program can be observed by the responses to events such as user actions, I/O signals, others through variations in the time the control process or messages from other programs. EDP is the preva- takes for dispatching their events. Figure 1 illustrates lent programming paradigm for graphical user interfaces, such a scenario for a loop that is shared between an at- web clients, and it is rapidly gaining importance for tacker A and a victim V. server-side and network programming. For instance, the Attacks based on observable contention of shared re- HTML5 standard [2] mandates that user agents be imple- sources have a long history [25] and an active present [8, mented using EDP, similarly, Node.js, memcached, and 27, 37]; however, attacks against shared event loops have Nginx, also rely on EDP. so far only been considered from a theoretical point of In EDP, each program has an event loop which con- view [22]. Here, we perform the first attacks against real sists of a FIFO queue and a control process (or thread) EDP-based systems. Specifically, we target shared event that listens to events. Events that arrive are pushed into loops in the two central processes of Google’s Chrome USENIX Association 26th USENIX Security Symposium 849 web browser: The host process, whose event loop is event loop are significantly less noisy or require less priv- shared between all requests for common resources, such ileges than from other channels [20, 38, 18]. as network and user interface; and the renderer pro- cesses, whose loops can be shared between Javascript tasks of different tabs or iframes. We build infrastructure that enables us to spy on both loops from a malicious HTML page. This is facilitated by the asynchronous programming model used in both Chrome and Javascript. Asynchronous function calls trigger new tasks that are appended to the same queue, in contrast to synchronous calls which are simply pushed onto the current task’s call stack and executed without preemption, blocking the loop. • For the event loop of the renderer we rely on the postMessage API, which is a Javascript feature for cross-window communication based on asyn- chronous callbacks. By posting messages to our- selves we can monitor the event loop with a resolu- tion of 25 µs, with only one task in the loop at each Figure 2: Delays observed while loading different web point in time. pages, by an attacker tab sharing the renderer process. • For the event loop of the host process we rely on Horizontal axis depicts elapsed real time, vertical axis two different mechanisms: network requests to non- depicts time taken by the event loop for processing the routable IP addresses, which enter the loop and tasks inserted by the attacker. All pages are clearly dis- abort very quickly, providing a resolution of 500 µs; tinguishable, both by the human eye and by classification and SharedWorkers, whose messages pass through techniques. the event loop of the host process, providing a reso- lution of 100 µs. 3. We demonstrate that shared event loops can We use the information obtained using these techniques be used to transmit information between cross-origin in three different attacks: pages.Specifically, we implement a covert channel with a bandwidth of 200 bit/s through the renderer’s main 1. We show how event delays during the loading thread event loop, and another one working cross- phase, corresponding to resource requests, parsing, processes of 5 bit/s. rendering and Javascript execution, can be used to uniquely identify a web page. Figure 2 visualizes this Our attacks show that event loops can be successfully effect using three representative web pages. While this spied on even with simple means. They work under the attack shares the goal with the Memento attack [21], assumption that event loops behave as FIFO queues; in the channels are quite different: First, in contrast to reality, however, Chrome’s event loop has a more so- Memento, we find that the relative ordering of events is phisticated structure, relying on multiple queues and a necessary for successful classification, which motivates policy-based scheduler. We believe that this structure can the use of dynamic time warping as a distance measure. be leveraged for much more powerful attacks in the fu- Second, we show that page identification through the ture. event loop requires only minimal training: we achieve recognition rates of up to 75% and 23% for the event 2 Isolation Policies and Sharing of Event loops of the renderer and host processes, respectively, Loops in Chrome for 500 main pages from Alexa’s Top sites. These rates are obtained using only one sample of each page for the In this section we revisit the same origin policy and its training phase. variants. We then discuss the relationship of these poli- cies with the Chrome architecture, where we put a special 2. We illustrate how user actions in cross-origin pages focus on the way in which event loops are shared. can be detected based on the delays they introduce in the event loop. In particular, we mount an attack against 2.1 Same Origin Policy Google OAuth login forms, in which we measure the time between keystrokes while the user is typing a pass- The Same-Origin Policy (SOP) is a central concept in word. The timing measurements we obtain from the the web security model: The policy restricts scripts on a 850 26th USENIX Security Symposium USENIX Association web page to access data from another page if their origins • the CompositorThread, which improves respon- differ. Two pages have the same origin if protocol, port siveness during the rendering phase by allowing the and host are equal. user to scroll and see animations while the main The demand for flexible cross-origin communication thread is busy, thanks to a snapshot of the page’s has triggered the introduction of features such as domain state. relaxation, the postMessage API, Cross-origin Resource Each of the threads in the host and renderer processes Sharing (CORS), Channel Messaging, Suborigins, or the maintains at least one event loop that is largely a FIFO Fetch API. This feature creep comes with an increase in queue. Inter-thread and inter-process communication are browser complexity and attack surface, which has mo- carried out via message passing through these queues. tivated browser vendors to move towards more robust We next discuss scenarios where pages of different origin multi-process architectures. can share the event loops of host and renderer processes. In Section 3 we show how this sharing can be exploited 2.2 Overview of the Chrome Architecture for eavesdropping. The Chrome architecture is segmented into different op- 2.3 Sharing in the Renderer Processes erating system processes. The rationale for this seg- mentation is twofold: to isolate web content from the Chrome supports different policies that govern how web host [6], and to support the enforcement of origin poli- applications are mapped to renderer processes, and that cies by means of the OS [30]. For achieving this seg- influence whether or not event loops are shared. mentation, Chrome relies on two processes: The default policy is called process-per-site- instance. It requires using a dedicated renderer HOST PROCESS process for each instance of a site. Here, a Main Thread site is defined as a registered domain plus a I/O Thread scheme. For example, https://docs.google.com and https://mail.google.com:8080 are from the same site – but not from the same origin, as they differ in subdomain RENDERER A RENDERER B and port. A site instance is a collection of pages from MainThread MainThread the same site that can obtain references to each other IOChildThread IOChildThread CompositorThread CompositorThread (e.g., one page opened the other in a new window using Javascript). Figure 3: Overview of Chrome’s architecture. The other supported policies are more permissive. For example, the process-per-site policy groups all instances of a site in the same renderer process, trading robust- The host process runs the top-level browser window. ness for a lower memory overhead. The process-per-tab It has access to system resources such as network, file policy dedicates one renderer process to each group of system, UI events, etc., which it manages on behalf of script-connected tabs. Finally, the single-process policy the unprivileged renderer processes. The host process lets both the host and renderer run within a single OS runs several threads; the most relevant ones are: process (only used for debugging purposes). • the CrBrowserMain thread, which handles, e.g., Even in the restrictive default process-per-site- user interaction events, and instance policy, there are some situations that force • the IOThread, which handles, e.g., IPC, network Chrome to host documents from different sites in the stack, and file system. same renderer process, causing them to share the event The renderer processes are sandboxed processes re- loop: sponsible for parsing, rendering and Javascript execu- • Iframes are currently hosted in the same process as tion. Communication with the host process is done via their parent. an inter-process communication (IPC) system based on • Renderer-initiated navigations such as link clicks, message passing. Each renderer runs several threads; the form submissions, and scripted redirections will most relevant ones are: reuse the same renderer as the origin page. • When the number of renderer processes exceeds a • the MainThread where resource parsing, style cal- certain threshold, Chrome starts to reuse existing culation, layout, painting and non-worker Javascript renderers instead of creating new ones. runs, • the IOChildThread, which handles IPC communi- On (64-bit) OSX and Linux, the threshold for reusing cation with the host process, and renderers is calculated by splitting half of the physical USENIX Association 26th USENIX Security Symposium 851 RAM among the renderers, under the assumption that 3 Eavesdropping on Event Loops in each consumes 60MB.1 In our experiments, on a ma- Chrome chine with 4 GB of RAM we could spawn 31 new tabs before any renderer was shared, whereas on a machine In this section we describe how to violate the SOP by with 8 GB of RAM we observed a threshold of approx- eavesdropping on the event loops of Chrome’s host and imately 70 renderers. There is no apparent grouping renderer processes. For each of these processes, we de- policy for the pages that can share a process when this scribe potential threat scenarios and present a simple threshold is exceeded, except for tabs in Incognito mode HTML page executing Javascript that can be used for not being mixed up with “normal” tabs. In particular, spying. We then present our monitoring tool to visual- we do not observe any preference for similar origins, ize the event loops of the browser. same sites, or secure versus insecure pages. In fact, even filesystem pages (loaded with file://) can co-reside 3.1 The Renderer Process Event Loop with an arbitrary HTTP site. 3.1.1 Threat Scenarios There are several scenarios in which an adversary site A can share the event loop of the renderer’s main thread 2.4 Sharing in the Host Process with a victim site V. These scenarios are based on Chrome’s policy for mapping sites to renderers, see Sec- The Chrome sandbox restricts access of renderers to tion 2.3. We give two examples: privileged actions. In particular, renderers have to com- • Malicious advertisement. In this scenario, A runs municate with the host process for network requests or as an advertisement iframed in V. The SOP pro- user input. The corresponding messages of all render- tects V’s privacy and itegrity by logically isolating ers pass through the event loop of the host process’ I/O both execution environments. However, A’s iframe thread. is able to execute Javascript on V’s event loop, en- We illustrate this communication using two different abling it to gather information about the user behav- examples: how user actions flow from the host to the cor- ior in V. responding renderer process, and conversely, how net- • Keylogger. In this scenario, A pops up a login form work requests flow from a renderer to the host process. to authenticate its users via V’s OAuth. Because the operation does not ask for special privileges and the • UI flow: User actions such as mouse movements password is never sent to A, the victim could trust or clicks enter the browser through the main thread it and fill the form. Meanwhile, A’s page monitors of the host process. The host main thread commu- keystroke timings (see Section 4.2), which can be nicates the user event to the corresponding renderer used for recovering user passwords [32]. by message passing between their I/O event loops, and the render acknowledges the receipt of this mes- 3.1.2 Monitoring Techniques sage. Even events with no Javascript listeners oc- cupy the event loop of the renderer’s main thread To monitor the renderer’s event loop it is sufficient to for a measurable interval. continuously post asynchronous tasks and measure the • Net stack: Chrome’s net stack is a complex cross- time interval between subsequent pairs of events. We platform network abstraction. Any network request measure the monitoring resolution in terms of the inter- by a renderer is passed to the I/O thread of the host val between two subsequent measurement events on an process, which forwards it to a global resource dis- otherwise empty loop. patcher that will pass it to a worker to fulfill the The most common way of posting asynchronous tasks request. This worker will open a connection, if programmatically in Javascript is setTimeout. How- necessary, and request the resource. After the re- ever, the resolution can be more than 1000 ms for inac- quest is done, the response headers are received and tive tabs, rendering this approach useless for the purpose sent back to the renderer process, which will re- of spying. To increase the resolution, we instead use the spond with an ACK after reading, Finally, the body postMessage API for sending asynchronous messages is received and the corresponding callbacks are trig- to ourselves. gered. The code in Listing 1 shows how this is achieved. The call to performance.now() in line 2 of the function loop returns a high-resolution timestamp 1 On Android there is no threshold since the OS suspends idle pro- that is saved as described below. The call to cesses. self.postmessage(0,’*’) in line 3 posts message 852 26th USENIX Security Symposium USENIX Association 1 function loop () { While all of these features reduce the effectiveness of our 2 save ( performance . now () ) attacks, it is interesting to think of them as potential side- 3 self . postMessage (0 , '* ') channels by themselves. For example, observable GC 4 } 5 self . onmessage = loop and JIT events can reveal information about a program’s 6 loop () memory and code usage patterns, respectively [29]. Listing 1: Javascript code to monitor the main thread’s event loop with the postMessage API. 3.2 The Host Process Event Loop 3.2.1 Threat Scenarios “0” into the renderer’s event loop, where the second argu- The Chrome sandbox ensures that all of the renderer’s ment “*” indicates no restriction on the receiver’s origin. network and user interaction events pass through the host Line 5 registers the function loop as an event listener, process’ I/O event loop, see Section 2.4. We describe two which enables it to receive the messages it has posted. threat scenarios where this could be exploited. This causes loop to recursively post tasks, while keep- • Covert channel. Pages of different origins running ing the render responsive since other events are still being in different (disconnected) tabs can use the shared processed. event loop to implement a covert channel, violat- In order to minimize the noise introduced by the ing the browser’s isolation mechanisms. This will measurement script itself, the function save in line 2 work even if one (or both) pages run in incognito uses a pre-allocated typed array (Float64Array) to mode. This channel can be used for tracking users store all the timing measurements. Contrary to normal across sessions, or to exfiltrate information from Javascript’s sparse arrays, typed arrays avoid memory re- suspicious web pages without network traffic. allocations and thus noisy garbage collection rounds, see • Fingerprinting. A tab running a rogue page of A can below. With that we achieve an average delay between identify which pages are being visited by the user in two consecutive tasks of around 25 µs on our target ma- other tabs by spying on the shared event loop. De- chine. This resolution is sufficient to identify even short tecting the start of a navigation is facilitated by the events. For example, a single mouse movement event fact that the I/O thread blocks for a moment when (without explicit event listener) consumes around 100 µs. the user types in a URL and presses enter. 3.2.2 Monitoring Techniques 3.1.3 Interferences There are many ways to post asynchronous tasks into the In modern browsers there are several sources of noise event loop of the host process; they differ in terms of the that affect measurement precision, beside the obvious ef- resolution with which they enable monitoring the event fect of the underlying hardware platform and OS. They loop and the overhead they imply. Below we describe include: two of the techniques we used. • Just-in-time compilation (JIT). JIT can trigger code optimization or deoptimization, in the case of Network Requests. The first technique is to use net- Chrome by the CrankShaft and Turbofan compil- work requests to systematically monitor the event loop ers, at points in time that are hard to predict. For of the I/O thread of the host process. A valid network re- our measurements we rely on a warm-up phase of quest may take seconds to complete, with only the start about 150 ms to obtain fully optimized code. and end operations visible in the loop, which provides • Garbage collection (GC). In the case of V8, GC in- insufficient resolution for monitoring. cludes small collections (so-called scavenges) and To increase the resolution, we make use of non- major collections. Scavenges are periodical and fast routable IP addresses. The corresponding requests en- (< 1 ms); but major collections may take > 100 ms, ter the I/O thread’s event loop, are identified as invalid distributed into incremental steps. In our data, scav- within the browser, and trigger the callback without any enges are easily identifiable due to their periodicity, DNS resolution or socket creation. This mechanism pro- while major collections could be spotted due to their vides a monitoring resolution of 500 µs and has the addi- characteristic size. On some browsers, such as Mi- tional benefit of being independent from network noise. crosoft’s Internet Explorer, GC rounds can be trig- Listing 2 shows the code of our monitoring proce- gered programmatically, which helps to eliminate dure. We rely on the Javascript Fetch API for posting the noise from the measurements enabling more precise network requests. The Fetch API provides an interface attacks [11]. for fetching resources using promises, which are ideal to USENIX Association 26th USENIX Security Symposium 853 manage asynchronous computations thanks to their sim- 3.2.3 Interferences ple syntax for handling callbacks. In line 2 we request and save a high-resolution timestamp. In line 3 we re- There are many different sources of noise and uncertainty quest a non-routable IP address, and set the rejection call- in the I/O thread of the host process. The most notable back of the promise to self, to recursively run when the ones include the interleaving with the host’s main thread request fails. and the messages from other renderers, but also the GPU process and browser plugins. While these interferences could potentially be exploited as side channels, the noise 1 function loop () { 2 save ( performance . now () ) becomes quickly prohibitive as the loop gets crowded. 3 fetch ( new Request ( ' http ://0/ ') ) . 4 catch ( loop ) 5 } 3.3 The LoopScan Tool 6 loop () We implement the eavesdropping techniques described Listing 2: Javascript code to monitor the host’s I/O in Sections 3.1 and 3.2 in a tool called LoopScan, which thread using network requests. enables us to explore the characteristics of the side chan- nel caused by sharing event loops. LoopScan is based on a simple HTML page that monitors the event loops Shared Workers. The second technique relies on web of the host and renderer processes. It relies on the D3.js workers, which is a mechanism for executing Javascript framework, and provides interactive visualizations with in the background. Web workers that are shared between minimap, zooming, and scrolling capabilities, which fa- multiple pages are usually implemented in a dedicated cilitates the inspection of traces. For example, Figure 8 OS process; this means they communicate via IPC and, is based on a screenshot from LoopScan. therefore, can be used to spy on the I/O thread of the host LoopScan’s functionality is in principle covered process. This mechanism provides a monitoring resolu- by the powerful Chrome Trace Event Profiling Tool tion of 100 µs. Listing 3 shows the code of our worker- (about:tracing) [3], which provides detailed flame graphs for all processes and threads. However, LoopScan 1 onconnect = function reply ( e ) { has the advantage of delivering more accurate timing 2 let port = e . ports [0] information about event-delay traces than the profiler, 3 port . onmessage = function () { since loading a page with the Trace Event Profiling tool 4 port . postMessage (0) severely distorts the measurements. LoopScan source is 5 } 6 } publicly available at https://github.com/cgvwzq/ loopscan. 1 const w = new SharedWorker ( ' pong . js ') 2 function loop () { 3 save ( performance . now () ) 4 Attacks 4 w . port . postMessage (0) 5 } In this section we systematically analyze the side channel 6 w . port . onmessage = loop 7 loop () caused by sharing event loops in three kinds of attacks: a page identification attack, an attack where we eavesdrop Listing 3: Javascript code to monitor the host’s on user actions, and a covert channel attack. For all at- I/O thread using SharedWorkers. The first snippet tacks we spy on the event loops of the renderer and the is the worker’s ‘pong.js’ file. Second snippet is host processes, as described in Sections 3.1 and 3.2. We the Javascript code that monitors the I/O thread by performed these attacks over the course of a year, always communicating with the worker. using the latest stable version of Chrome (ranging from v52-v58). The results we obtain are largely stable across based monitoring procedure. The first snippet defines the the different versions. worker’s job, which consists in replying to each received message. In the second snippet, we register the worker in 4.1 Page identification line 1. In lines 2-7 we record a timestamp and recursively send messages to the worker, analogous to Listing 1. As We describe how the event-delay trace obtained from a result, we measure the round-trip time from the page to spying on event loops can be used for identifying web- the worker, which reflects the congestion in the I/O event pages loaded in other tabs. We begin by explaining loop. Note that one can further increase the measure- our data selection and harvesting process and the cho- ment resolution by recording the time in each endpoint sen analysis methods, then we describe our experimental and merging the results. setup and the results we obtain. 854 26th USENIX Security Symposium USENIX Association 4.1.1 Sample Selection We measure the timing on a Chrome instance with two tabs, one for the spy process and the other for the target We start with the list of Alexa Top 1000 sites, from page. For the renderer process, we gather data on all which we remove duplicates. Here, duplicates are sites machines; for the host process on (2) and (3). Overall, that share the subdomain but not the top-level domains we thus obtain five corpora of 15.000 traces each. (e.g., “google.br” and “google.com”) and that are likely to have similar event-delay traces. From the remaining list, we randomly select 500 sites as our sample set. This 4.1.3 Classification reduction facilitates a rigorous exploration of the data and the parameter space. Event Delay Histograms. Our first approach is to cluster the observed event delays around k centers, and to transform each trace into a histogram that represents 4.1.2 Data Harvesting the number of events that fall into each of the k classes. We then use the Euclidean distance as a similarity mea- We visit each page in the sample set 30 times for both the sure on the k-dimensional signatures. renderer and the host process, to record traces of event- delays during the loading phase. This approach is inspired by the notion of memprints in [21]. It appears to be suitable for classifying event- The event-delay traces for the renderer process con- delay traces obtained from event loops because, for ex- sist of 200.000 data items each. On our testing machine, ample, static pages with few external resources are more the measurement resolution (i.e. the delay between two likely to produce long events at the beginning and stabi- subsequent measurement events on an otherwise empty lize soon, whereas pages with Javascript resources and loop) lies at approximately 25 µs. That is, each trace animations are likely to lead to more irregular patterns captures around 5 seconds (200.000·25 µs = 5 s) of the and produce a larger number of long delays. Unfortu- loading process of a page in the sample set. nately, our experimental results were discouraging, with The event-delay traces for the host process consist of less than a 15% of recognition rate in small datasets. 100.000 data items each. The measurement resolution lies in the range of 80 − 100 µs, i.e. each trace captures around 9 s of the loading process of a page. Dynamic Time Warping. Our second approach is We automate the harvesting procedure for the renderer to maintain temporal information about the observed process as follows: events. However, the exact moments at which events 1. Open a new tab via occur are prone to environmental noise. For example, target = window.open(URL, ’_blank’); 2 network delays will influence the duration of network 2. Monitor the event loop until the trace buffer is full requests and therefore the arrival of events to the event 3. Close the tab loop. Instead, we focus on the relative ordering of events 4. Send the trace to the server as a more robust feature for page identification. 5. Wait 5 seconds and go to 1 with next URL This motivates the use of dynamic time warping The harvesting procedure for the host process differs (DTW) [7] as a similarity measure on event-delay traces. only in that we use the rel="noopener" attribute in or- DTW is widely used for classifying time series, i.e. se- der to spawn a new renderer. quences of data points taken at successive and equally We conducted measurements on the following three spaced points in time. DTW represents a notion of dis- machines: tance that considers as “close” time-dependent data of similar shape but different speed, i.e. DTW is robust to 1. Debian 8.6 with kernel 3.16.0-4-amd64, running on horizontal compressions and stretches. This is useful, an Intel i5 @ 3.30GHz x 4 with 4 GB of RAM, and for example, when one is willing to assign a low distance Chromium v53; score to the time series “abc“ and “abbbbc‘, insensitive to the prolonged duration of “b“. Formally, DTW com- 2. Debian 8.7 with kernel 3.16.0-4-amd64, running on pares two time series: a query, X = (x1 , ..., xn ), and a ref- an Intel i5-6500 @ 3.20GHz x 4 with 16 GB of erence, Y = (y1 , ..., ym ). For that we use a non-negative RAM, and Chromium v57; and distance function f (xi , yi ) defined between any pair of el- ements xi and y j . The goal of DTW is to find a matching 3. OSX running on a Macbook Pro 5.5 with In- of points in X with points in Y , such that (1) every point tel Core 2 Duo @ 2.53GHz with 4 GB of RAM, is matched, (2) the relative ordering of points in each se- and Chrome v54. quence is preserved (monotonicity), (3) and the cummu- 2 Note that this requires disabling Chrome’s popup blocker from lative distance (i.e. the sum of the values of f ) over all “chrome://settings/content”. matching points is minimized. This matching is called a USENIX Association 26th USENIX Security Symposium 855 warping path, and the corresponding distance is the time warping distance d(X,Y ). Figure 5: The top figure represents a raw trace of 200.000 time measurements from the renderer’s main thread ex- tracted while loading “google.com”. The bottom figure displays the same data after being converted into a time series with P = 20 ms, i.e. using only 250 data points. Figure 4: The path in the upper right square represents The difference in the height of the peaks is due to the ac- the optimal alignment between points in the time se- cumulation of small events in the raw data, which are not ries corresponding to ’google.com’ (horizontal axis) with perceptible in the top figure. points in the time series of ’youtube.com’ (vertical axis). Figure 4 visualizes a warping path between the • Windowing is a heuristic that enforces a global con- time series corresponding to event-delay traces observed straint on the envelope of the warping path. It speeds while loading different webpages. up DTW but will not find optimal warping paths that lie outside of the envelope. Two well-established constraint 4.1.4 Speed-up Techniques regions are the Sakoe-Chiba band and the Itakura paral- lelogram, see Figure 6. Unfortunately, the time required for computing d(X,Y ) is quadratic in the length of the input sequences and does (a) (b) not scale up to the raw data obtained in our measure- ments. We rely on two kinds of speed-up techniques, one at the level of the data and the other at the level of the algorithm: At the level of data, we reduce the dimension of our data by applying a basic sampling algorithm: We split the raw trace into groups of measurements corresponding to time intervals of duration P, and replace each of those groups by one representative. This representative can be Figure 6: A global window constraint defines an enve- computed by summing over the group, or by taking its lope limiting the search space for optimal warping paths: average, maximum or minimum. The sum function gen- (a) Itakura parallelogram, and (b) Sakoe-Chiba band. erally yields the best results among different sampling functions and is the one that we use onwards. Sampling reduces the size of the traces by a factor of P/t, where t • Step patterns are a heuristic that puts a local con- is the average duration of an event delay. Figure 5 shows straint on the search for a warping path, in terms of re- two plots with the raw data taken from a renderer’s main strictions on its slope. In particular, we rely on three thread loop, and its corresponding time series obtained well-known step patterns available in R. Intuitively, the after sampling. symmetric1 pattern favors progress close to the diagonal, At the algorithmic level, we use two sets of tech- the symmetric2 pattern allows for arbitrary compressions niques for pruning the search for the optimal warping and expansions, and the asymmetric forces each point in path, namely windowing and step patterns [15]. the reference to be used only once. 856 26th USENIX Security Symposium USENIX Association Sakoechiba - symmetric1 Sakoechiba - asymmetric Itakura - symmetric1 100% 100% 100% 75% 75% 75% 50% 50% 50% P 5 TraceDuration 25% 25% 25% 10 4s 20 2s 0% 0% 0% 50 1s 1 5 10 30 50 100 1 5 10 30 50 100 1 WindowSize WindowSize WindowSize Figure 7: Web page identification performance after tuning with traces from the renderer on Linux machine (1). Effect of P, traceDuration, and windowSize, with three combinations of stepPattern and windowType. 4.1.5 Parameter tuning cesses on each individual machine, as well as through the renderer process across two different machines. The possible configurations of the techniques presented To this end, we select the top configuration for each in Section 4.1.4 create a large parameter space, see Ta- corpus from the tuning phase and carry out a 10-fold ble 1 for a summary. cross-validation. In each of the 10 rounds, we partition Parameter Values Description the validation set into a training set that contains one traceDuration 1000, 2000, 4000 Trace duration (ms) trace of each page, and a testing set that contains three P 5, 10, 20, 50 Sampling interval (ms) different (out of the 14 available) traces of each page. windowType itakura, sakoechiba Window constraint windowSize 1, 5, 10, 30, 50, 100 Window size For each of the traces in the testing set, we compute the symmetric1, symmetric2, set of k closest matches in the training set according to stepPattern Step pattern asymmetric the time warping distance. Table 1: List of parameters tuned for optimizing web We measure performance in terms of the k-match page identification rate, which is the percentage of pages in the testing set for which the true match is within the set of k closest We systematically identify the optimal parameter con- matches. We abbreviate the 1-match rate by recognition figuration for each event loop on each machine. To avoid rate, i.e. the percentage of pages where the best match is overfitting, we divide our dataset of 30 traces (per page, the correct one. The result of the cross-validation is the loop, and machine) into 15 traces for tuning and 15 for average k-match rate over all 10 rounds. cross-validation. For each parameter configuration we Table 2 summarizes our experiments. We highlight the perform a lightweight version (with 3 rounds) of the eval- following results: uation phase described in Section 4.1.6. Figure 7 visual- izes an extract of the results we obtain for the renderer k 1 3 5 10 process of the Linux (1) machine. The tuning phase 76.7 % 86.7 % 88.8 % 91.1 % (1) yields the following insights: Renderer sym1,sakoe, P = 5, windowSize = 100 • The optimal parameters depend on the loop but ap- 58.2 % 68.6 % 71.8 % 75.1 % pear to be stable across machines. Renderer sym1,sakoe, P = 5, windowSize = 100 (2) • Measuring the loading phase during 2 seconds is 16.2 % 23.2 % 27.9 % 36.1 % I/O host sufficient for recognition of a webpage; the gain in recog- sym1,sakoe, P = 20, windowSize = 30 nition from using longer traces is negligible. Renderer 61.8 % 74.5 % 78.4 % 83.1 % • P and windowSize are the parameters with the sym1,sakoe, P = 5, windowSize = 100 (3) biggest impact on the recognition rate. However, they 23.48 % 32.9 % 38.1 % 46.6 % I/O host sym1,sakoe, P = 20, windowSize = 30 also have the biggest impact on the computational cost (the optimal choice being most expensive one). Table 2: 10-fold cross-validation results on different ma- • The combination of stepPattern = symmetric1 and chines and different event loops, with the best config- windowType = sakoechiba generally yields the best re- uration after tuning. Machines (1) and (2) refer to the sults. Linux desktops, (3) to the OSX laptop, as described in Section 4.1.2. 4.1.6 Experimental Results We evaluate the performance of page identification • We can correctly identify a page by spying on the through the shared event loops of host and renderer pro- renderer from (1) in up to 76.7% of the cases, and cor- USENIX Association 26th USENIX Security Symposium 857 rectly narrow down to a set of 10 candidates in up to 4.2 Detecting User Behavior 91.1% of the cases. In this section we show that it is possible to detect user • We can correctly identify a page though the host actions performed in a cross-origin tab or iframe, when process from (3) in up to 23.48% of the cases, and nar- the renderer process is shared. We first describe an attack row down to a set of 10 candidates in up to 46.6% of the recovering the inter-keystroke timing information against cases. Google’s OAuth login forms, which provides higher pre- • We stress that these recognition rates are obtained cision than existing network-based attacks [32]. using a single trace for training. • Recognition is easier through the renderer than 4.2.1 Inter-keystroke Timing Attack on Google’s through the host. This is explained by the difference OAuth login form in noise and measurement resolution, see Section 3.2.3. Furthermore, most operations on the host only block the Many web applications use the OAuth protocol for user I/O thread while signaling their start and completion, authentication. OAuth allows users to login using their whereas the renderer is blocked during the entire exe- identity with trusted providers, such as Google, Face- cution of each Javascript task. book, Twitter, or Github. On the browser, this process is commonly implemented as follows: • We observe different recognition rates on different 1. A web application A pops up the login form of a machines. However the homogeneity in hardware and trusted provider T; software of Macbooks facilitate reuse of training data 2. User V types their (name and) password and sub- across machines, which may make remote page identi- mits the form to T; fication more feasible. 3. T generates an authorization token. • We obtain recognition rates below 5 % for recog- Because the window of the login form shares the event nition across machines (1) and (3). A reason for this loop with the opener’s renderer, a malicious A can eaves- poor performance is that events on the OSX laptop of- drop on the keystroke events issued by the login form. ten take 2x-5x more time than on the Linux desktop ma- chine. This difference is reflected in the height of the 10.00 peaks (rather than in their position), which is penalized 4.00 by DTW. Normalizing the measurements could improve 2.00 cross-machine recognition. 1.00 0.40 The code and datasets used for tuning and cross- 0.20 validation are available as an R library at https:// 0.10 0.06 github.com/cgvwzq/rlang-loophole. 0.04 0.02 19780.000 19785.000 19790.000 19795.000 19800.000 19805.000 Figure 8: Delay pattern generated by a keystroke in the Google OAuth login form, measured across origins on 4.1.7 Threats to Validity Chrome Canary v61 on OSX. The two consecutive de- lays of approx. 2ms each, correspond to keydown and We perform our experiments in a closed-world scenario keypress event listeners. with only 2 tabs (the spy and the victim) sharing an event loop. In real world scenarios there can be more pages Figure 8 depicts the event-delay trace of a keystroke concurrently running the browser, which will make de- as seen by an eavesdropper on the renderer’s event loop. tection harder. The worst case for monitoring the host The trace contains two characteristic consecutive delays, process occurs when a tab performs streaming, since the caused by the keydown and keypress event listeners. We loop gets completely flooded. The renderer’s loop, how- use this observation to identify keystrokes, by scanning ever, is in general more robust to noise caused by other the event-delay trace for pairs of consecutive delays that tabs in the browser. are within a pre-defined range, forgoing any training or On the other hand, our attacks do not make any use of offline work. Listing 4 contains the script that performs the pages’ source code or of details of Chrome’s schedul- this operation. We define 0.4 ms as a lower bound, and ing system with priority queues, the GC with periodic 3.0 ms as an upper bound for the range. We chose this scavenges, or the frame rendering tasks. We believe that threshold before gathering the data, by manual inspection taking into account this information can significantly im- of a few keystroke events. Note that this calibration could prove an adversary’s eavesdropping capabilities and en- be done automatically, based on the victim’s interactions able attacks even in noisy, open-world scenarios. with a page controlled by an attacker. 858 26th USENIX Security Symposium USENIX Association 1 const L = 0.4 , U = 3.0 , keys = [] 3. V and A send to the server the timestamps of the 2 real and the detected keystrokes, respectively. 3 for ( let i =1; i < trace . length -1; i ++) { 4. We compute the accuracy of the detected 4 let d1 = trace [ i ] - trace [i -1] , 5 d2 = trace [ i +1] - trace [ i ] keystrokes, where we take the timestamps of 6 the real keystrokes as ground truth. Matching the 7 if (L < d1 < U && L < d1 < U ) { timestamps requires taking into account the delay 8 keys . push ( trace [ i ]) 9 } (6 − 12 ms on our machine) between Selenium 10 } triggering an event, and Chrome receiving it. We use as inter-keystroke timings random delays uni- Listing 4: Pseudo-Javascript code to detect formly drawn from 100 − 300 ms. This choice is inspired keystrokes in a trace of timestamps gathered by by [20], who report on an average inter-keystroke delay the code in Listing 1. We classify a timestamp as of 208 ms. Using random delays is sufficient for evalu- a keystroke if the differences to the previous and ating the accuracy of eavesdropping on keystrokes, but subsequent timestamps (d1 and d2) are both in a it obviously does not reveal any information about the predefined range. password besides its length. 4.2.3 Experimental Results 4.2.2 Experimental Evaluation We perform experiments with 10.000 passwords ex- To evaluate the effectiveness of this attack, we have tracted from the RockYou dataset, where we obtain the implemented a malicious application A that extracts following results: the inter-keystroke timing information from a user V logging-in via Google’s OAuth. The focus of our evalu- • In 91.5% of the cases, our attack correctly identifies ation is to determine the accuracy with which keystroke the length of a password. 4 In 2.2% of the cases, the timings can be measured through the event loop. A full attack misses one or more characters, and in 6.3% keystroke recovery attack is out of scope of this paper; of the cases it reports spurious characters. for this refer to [32]. • For the passwords whose length was correctly iden- tified, the average time difference between a true keystroke and a detected keystroke event is 6.3ms, which we attribute mostly to the influence of Se- lenium. This influence cancels out when we com- pute the average difference between a true inter- keystroke delay and a detected inter-keystroke de- lay, which amounts to 1.4 ms. The noise of these measurements is low: We observe a standard devia- tion of 6.1 ms, whereas the authors of [20] report on 48.1 ms for their network based measurements. Overall, our results demonstrate that shared event Figure 9: Experimental setup for evaluating effectiveness loops in Chrome enable much more precise recovery of of automatic, cross-renderer keystroke detection. keystroke timings than network-based attacks. More- over, this scenario facilitates to identify the time when We simulate an inter-keystroke timing attack in 4 keystroke events enter the loop (from popping-up to form steps, which are described below and illustrated in Fig- submission), which is considered to be a major obstacle ure 9. for inter-keystroke timing attacks on network traffic [20]. 1. A Selenium3 script acting as V navigates to A, clicks Keystroke timing attacks based on monitoring on the login button (which pops up Google’s OAuth procfs [38] or CPU caches [18] can extract more fine- login form), types a password, and submits the grained information about keystrokes, such as contain- form. ment in a specific subsets of keys. However, they require 2. Meanwhile, the attacker A monitors the main filesystem access or are more susceptible to noise, due thread’s event loop using the attack described in to the resource being shared among all processes in the Section 4.2.1. system. In contrast, our attack enables targeted eaves- 3 Selenium (http://www.seleniumhq.org/) is a cross-platform dropping without specific privileges. testing framework for web applications that provides capabilities for 4 We configured Selenium to atomically inject characters that would programmatically navigating to web pages and producing user input. require multiple keys to be pressed. USENIX Association 26th USENIX Security Symposium 859 4.2.4 Open Challenges for Recognizing User Events affected by the execution of the corresponding Javascript event listener. We conclude by discussing two open challenges for recognizing user events, namely the detection of user events beyond keystrokes and the detection of events in 4.3 Covert Channel the browser’s host process. In this section we show how shared event loops in Chrome can be abused for implementing covert chan- Detecting User Events beyond Keystrokes A contin- nels, i.e. channels for illicit communication across ori- uous mouse movement results in a sequence of events, gins. We first consider the case of cross-origin pages each of which carrying information about the coordinates sharing the event loop of a renderer’s main thread be- of the cursor’s trajectory. These events are issued with an fore we turn to the case of cross-origin pages sharing the inter-event delay of 8 ms, and the (empty) event listener event loop of the host processes’ I/O thread. operation blocks the loop for approx 0.1 ms. The partic- ular frequency and duration of these events makes mouse 4.3.1 Renderer Process movements (or similar actions, like scrolling) easy to We implement a communication channel to transmit spot with LoopScan, as seen in Figure 10. messages from a sender page S to a cross-origin receiver page R running in the same renderer process. For this, we use a simple, unidirectional transmission scheme without error correction. Specifically, we encode each bit using a time interval of fixed duration tb . The op- timal configuration of tb depends on the system. In our Figure 10: Mouse movement captured by LoopScan tool. experiments we tried different values, with tb = 5 ms giv- The graph shows 3 delays of 0.1 ms duration (at t equals ing good results on different platforms: Chromium 52.0 3350, 3358 and 3366), with an inter-event delay of 8 ms. on Debian 64-bit and Chrome 53 on OSX. In each of those intervals we do the following: Likewise, mouse click events, corresponding to “up” • the sender S idles for transmitting a 0; it executes a or “down”, can be identified using LoopScan. Their blocking task of duration tˆ < tb for transmitting a 1. shape depends on the specific event listener of the spied • the receiver R monitors the event loop of the ren- web page and the HTML element being clicked. We ex- derer’s main thread using the techniques described pect that events with specific listeners are more easily in Section 3.1; it decodes a 0 if the length of the ob- detectable than events without registered event listeners, served tasks is below a threshold (related to tˆ), and that is, user actions that do not trigger Javascript exe- a 1 otherwise. cution. However, we can use the context in which the Transmission starts with S sending a 1, which is used by event occurs to reduce the search space. For instance, the agents to synchronize their clocks and start count- most mouse clicks only appear between two sequences ing time intervals. Transmission ends with S sending a of mouse movement events. null byte. With this basic scheme we achieve rates of We are currently investigating techniques that enable 200 bit/s. These numbers can likely be significantly the automatic identification of such patterns in event- improved by using more sophisticated coding schemes delay streams. A promising starting point for this are with error correction mechanisms; here, we are only in- existing on-line variants of dynamic time-warping [31]. terested in the proof-of-concept. We note that there are a number of alternative covert channels for transmitting information between Detecting User Events in the Host Process Our dis- pages running in the same renderer [1], e.g., us- cussion so far has centered on detecting user events in ing window.name, location.hash, history.length, the event loop of the renderer process. However, all user scrollbar’s position or window.frames.length. What events originate in the main thread of the host process distinguishes the event-loop based channel is that it does and are sent towards a specific renderer through the event not require the sender and receiver to be connected, i.e. loop of the host’s I/O thread. Hence, any user action can they do not need to hold references to each other in order in principle be detected by spying on the host. to communicate. Unfortunately, our current methods are not precise enough for this task, since the host’s I/O thread is more 4.3.2 Host Process noisy than the renderer’s main thread and the effect of a user action on the host process is limited to a short sig- We also implement a communication channel to transmit naling message, whereas the renderer’s main thread is messages between two cooperative renderer processes 860 26th USENIX Security Symposium USENIX Association sharing the host process. Transmission is unidirectional 5 Discussion from sender S to receiver R. Figure 11 visualizes how this channel can be used, even if one of the parties browses We have shown how sharing event loops leads to timing in Incognito mode. side-channels and presented different attacks on Chrome. We communicated our findings to the Chromium security team, who decided not to take action for the time being. Nevertheless, our results point to fundamental security issues in the event-driven architecture of browsers that eventually need to be addressed in a fundamental man- ner. Below, we discuss how other platforms are affected and present possible countermeasures. 5.1 Beyond Chrome We focus on Chrome in our analysis because it is the most widely used browser, and because it was the first one to implement a multi-process architecture. However, there are good reasons to expect similar side channels in Figure 11: Covert channel through the I/O event loop other browsers, as they all follow the same event-driven of the Chrome’s host process. Tabs in different renderer paradigm and rely on similar architectures. processes (one of them navigating in Incognito mode) For instance, recent Firefox versions with multi- communicate. process support5 also rely on a privileged browser pro- cess and multiple content processes that, unlike render- ers in Chrome, act as a pool of threads for each different origin (each with its own message queue). Despite this As before, we encode each bit using a time intervals difference, tests with LoopScan on Firefox version 55 of fixed duration tb . During each intervals we do the fol- show that congestion on both event loops is observable lowing: across origins and tabs. • the sender S idles for transmitting a 0; it posts N Specifically, we applied the monitoring technique for fetch requests into the I/O thread’s queue for send- the renderers described in Section 3.1.2 on a micro- ing a 1. benchmark with a set of 30 pages with 15 traces each. • the receiver R monitors the event loop of the I/O We achieved a recognition rate of 49%, which is be- thread of the host process using the techniques de- low the recognition rate achieved on Chrome for a set of scribed in Section 3.2. It decodes a 0 if the number 500 pages. A fair comparison between both architectures of observed events during time interval tb is below will require a better understanding of Firefox’s policy for a threshold, and 1 otherwise. mapping sites to threads and events to loops. The optimal values of N and tb highly depend on the ma- chine. In our experiments we achieve good results, work- 5.2 Countermeasures ing on different systems, with a tb = 200 ms and N = 350, The attacks presented in this paper rely on two capabili- which give us a 5 bit/s transmission rate. This rate is sig- ties of the adversary: (1) the ability to post tasks into the nificantly lower than for communication using the ren- loop’s queue with high frequency, and (2) the ability to derer event loop, which is explained by the difference in accurately measure the corresponding time differences. noise and monitoring resolution of both channels, as dis- cussed in Section 3.2.3. Rate Limiting. An obvious approach to counter (1) The threat scenario of this covert channel is more is to impose a limit on the rate at which tasks can be relevant than the previous one for the renderer loop. posted into an event loop. Unfortunately, rate limiting For example it could be used for exfiltrating informa- implies penalties on performance, which is especially tion from an attacked domain (on a tab executing mali- problematic for asynchronous code. cious Javascript). Using Workers (which are background At the level of the renderer, one possibility is to rely threads that run independently of the user interface) we on an accumulate and serve policy [22]. With this pol- can transfer information across origins, without affect- icy, the event loop accumulates all the incoming jobs ing the user experience and without generating network traffic. 5 Firefox’s Electrolysis (or e10s) project USENIX Association 26th USENIX Security Symposium 861 in a buffer for a period T , and then process and serves 6 Related Work all the accumulated jobs from party A, followed by all the jobs from V. This has the advantage of limiting the amount of information leaked while retaining high amor- Timing attacks on web browsers date back to Felten and tized throughput. Schneider [13], who use the browser cache to obtain in- At the level of the host process, where resource fetch- formation about a user’s browsing history. ing is one of the main performance concerns, setting any More recently, so-called cross-site timing attacks [10, bound on the processing rate is not acceptable. Here, it 35] have exploited the fact that the browser attaches seems more reasonable to monitor the IPC activity of all cookies to all requests, even when they are performed renderers and penalize or flag those who exhibit a bad or across origins. The presence or absence of these cookies anomalous behavior, e.g., along the lines of [39]. can be determined by timing measurements, which re- veals information about the user’s state on arbitrary sites. Reduce Clock Resolution. An obvious approach to A special case are cross-site search attacks [14], which counter (2) is to limit the resolution of available clocks. circumvent the same-origin policy to extract sensitive in- This has already been applied by browser vendors for formation, by measuring the time it takes for the browser mitigating other kinds timing channels, but these ef- to receive responses to search queries. forts are unlikely to succeed, as shown in [23]: Modern Other classes of browser-based timing attacks exploit browsers have a considerable number of methods to mea- timing differences in rendering operations [24, 33, 5], or sure time without any explicit clock. For instance, some simply use the browser as an entry point for Javascript recent exploits [16] use high-resolution timers build on that exploits timing channels of underlying hardware, for top of SharedArrayBuffers. The current resolution of example caches [26, 16], DRAM buffers [17], or CPU performance.now is limited to 5 µs, which makes mi- contention [9]. croarchitectural timing attacks difficult, but does not pre- clude the detection of Javascript events. Of those approaches, [9] is related to our work in that it identifies web pages across browser tabs, based on timing of Javascript and a classifier using dynamic time Full Isolation. As discussed in Section 2.2, Chrome’s warping. However, because the attack relies on CPU multi-process architecture tries to use a different ren- contention as a channel, it requires putting heavy load on derer for different origins, except for some corner all cores for monitoring. In contrast, our attack exploits cases. The “Site Isolation Project” is an ongoing ef- the browser’s event loop as a channel, which can be mon- fort to ensure a complete process-per-site-instance pol- itored by enqueing one event at a time. This makes our icy, that means: providing cross-process navigations, attack stealthy and more independent of the execution cross-process Javascript interactions and out-of-process platform. iframes. All this without inducing too much overhead. One open question is how to handle the system’s pro- To the best of our knowledge, we are first to mount cess limit, namely which sites should have isolation pref- side-channel attacks that exploit the event-driven archi- erence, or which heuristic for process reuse should be tecture of web browsers. Our work is inspired by a proof- used. A recent proposal, “IsolateMe” [4], puts the devel- of-concept attack [36] that steals a secret from a cross- opers in charge of requesting to be isolated from other origin web application by using the single-threadedness web content (even if it does not provide a firm guaran- of Javascript. We identify Chrome’s event-driven archi- tee). tecture as the root cause of this attack, and we show how this observation generalizes, in three different at- tacks against two different event loops in Chrome. CPU Throttling. Chrome v55 introduces an API that allows to limit how much CPU a background page is Finally, a central difference between classical site fin- allowed to use, and to throttle tasks when they exceed gerprinting [28, 19, 34, 12] approaches and our page this limit. This affects background tabs trying to spy identification attack is the adversary model: First, our ad- on the renderer’s main thread, but still allows spying versary only requires its page to be opened in the victim’s on (and from) any iframe and popup, as well as on the browser. Second, instead of traffic patterns in the vic- I/O thread of the host process through shared Workers. tim’s network, our adversary observes only time delays Moreover, background tabs with audio activity are not in the event queues of the victim’s browser. We believe affected, as they are always marked as foreground. Since that our preliminary results, with up to 76% of recogni- Chrome v57 pages (or tabs) are only subjected to throt- tion rate using one single sample for training in a closed- tling after 10 seconds in the background, which is too world with 500 pages, can be significantly improved by long to prevent the attacks in this paper. developing domain-specific classification techniques. 862 26th USENIX Security Symposium USENIX Association 7 Conclusions [15] G IORGINO , T. Computing and visualizing dynamic time warping alignments in r: The dtw package. JSS 31, 7 (2009), 1–24. In this paper we demonstrate that shared event loops in [16] G RAS , B., R AZAVI , K., B OSMAN , E., B OS , H., AND G IUF - Chrome are vulnerable to side-channel attacks, where a FRIDA , C. ASLR on the Line: Practical Cache Attacks on the MMU. In NDSS (2017), The Internet Society. spy process monitors the loop usage pattern of other pro- cesses by enqueueing tasks and measuring the time it [17] G RUSS , D., M AURICE , C., AND M ANGARD , S. Rowhammer.js: A remote software-induced fault attack in javascript. In DIMVA takes for them to be dispatched. We systematically study (2016), Springer. how this channel can be used for different purposes, such [18] G RUSS , D., S PREITZER , R., AND M ANGARD , S. Cache tem- as web page identification, user behavior detection, and plate attacks: Automating attacks on inclusive last-level caches. covert communication. In USENIX Security (2015), USENIX Association. 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