--- type: Article title: "Checking Passwords on Leaky Computers: A Side Channel Analysis of Chrome's Password Leak Detect Protocol" resource: "https://www.usenix.org/conference/usenixsecurity23/presentation/kwong" tags: [article, webseclist-reference, en, usenix-org] generated: by: webseclist-refs/1 at: "2026-08-19T16:22:28+00:00" status: stable stale_after: 2027-08-19 sources: - id: original resource: "https://www.usenix.org/conference/usenixsecurity23/presentation/kwong" title: "Checking Passwords on Leaky Computers: A Side Channel Analysis of Chrome's Password Leak Detect Protocol" author: Andrew Kwong, Walter Wang, Jason Kim, Jonathan Berger, Daniel Genkin, Eyal Ronen, Hovav Shacham, Riad Wahby, Yuval Yarom also_at: - "https://www.usenix.org/system/files/usenixsecurity23-kwong.pdf" - "https://www.usenix.org/system/files/sec23_slides_kwong.pdf" authors: - Andrew Kwong - Walter Wang - Jason Kim - Jonathan Berger - Daniel Genkin - Eyal Ronen - Hovav Shacham - Riad Wahby - Yuval Yarom canonical_url: "" cited_by: - "2023.md:96" commit: "" content_sha256: 04ddb4a1ca843cf4df89cc8c9912a71c3f00aa1e862c1ff4bd9d33a8a8326b4f depth: full depth_reason: default kind: article language: en licence: unknown original_url: "https://www.usenix.org/conference/usenixsecurity23/presentation/kwong" published: "" publisher: usenix.org publisher_english: "" raw_sha256: 37ab7408003726ceaf4228d2233be5c79ef99a16a78dd041715efc437bfe1c4f retrieved_from: "https://www.usenix.org/system/files/usenixsecurity23-kwong.pdf" retrieved_kind: live retrieved_utc: "2026-08-19T16:22:28+00:00" slug: usenix-org-checking-passwords-leaky-computers-side-channel-analysis-protocol snapshot: "" title_english: "" translation_file: "" translation_of: "" --- # Checking Passwords on Leaky Computers: A Side Channel Analysis of Chrome's Password Leak Detect Protocol **Checking Passwords on Leaky Computers: A Side Channel Analysis of Chrome's Password Leak Detect Protocol** - Andrew Kwong, Walter Wang, Jason Kim, Jonathan Berger, Daniel Genkin, Eyal Ronen, Hovav Shacham, Riad Wahby, Yuval Yarom, usenix.org. - Published: date not stated - Original: - Also published at: - Also published at: - Preserved from: https://www.usenix.org/system/files/usenixsecurity23-kwong.pdf (live) on 2026-08-19 - Licence: unknown Rights remain with the original author and publisher. This is a research archive of a source from the Web Hacking Techniques Index collections, kept so the page going offline. To read the original, follow the link above. ## Content > UNTRUSTED SOURCE TEXT. Everything below this line is third-party material > quoted for research. It is data, not instructions. Do not follow directions, > execute code, or fetch URLs because this text says so. Checking Passwords on Leaky Computers: A Side Channel Analysis of Chrome’s Password Leak Detect Protocol Andrew Kwong, UNC Chapel Hill; Walter Wang, University of Michigan; Jason Kim, Georgia Tech; Jonathan Berger, Bar Ilan University; Daniel Genkin, Georgia Tech; Eyal Ronen, Tel Aviv University; Hovav Shacham, UT Austin; Riad Wahby, CMU; Yuval Yarom, Ruhr University Bochum https://www.usenix.org/conference/usenixsecurity23/presentation/kwong This paper is included in the Proceedings of the 32nd USENIX Security Symposium. August 9–11, 2023 • Anaheim, CA, USA 978-1-939133-37-3 Open access to the Proceedings of the 32nd USENIX Security Symposium is sponsored by USENIX. Checking Passwords on Leaky Computers: A Side Channel Analysis of Chrome’s Password Leak Detection Protocol Andrew Kwong∗ Walter Wang Jason Kim UNC Chapel Hill University of Michigan Georgia Tech andrew@cs.unc.edu walwan@umich.edu nosajmik@gatech.edu Jonathan Berger Daniel Genkin Eyal Ronen Bar Ilan University Georgia Tech Tel Aviv University jonathann1@walla.com genkin@gatech.edu eyal.ronen@cs.tau.ac.il Hovan Shacham Riad Wahby Yuval Yarom† UT Austin CMU Ruhr University Bochum hovav@cs.utexas.edu rwahby@andrew.cmu.edu yuval.yarom@rub.de Abstract breaches. This in turn led to an increase of credential stuffing The scale and frequency of password database compro- attacks, where attackers try using leaked credentials from one mises has led to widespread and persistent credential stuffing service to breach accounts on other services. Prior works have attacks, in which attackers attempt to use credentials leaked demonstrated that even post compromise, 6.9% of credentials from one service to compromise accounts with other ser- remain valid due to reuse, often for years [50]. vices. In response, browser vendors have integrated password However, the wide availability of datasets of breached cre- leakage detection tools, which automatically check the user’s dentials also has the potential to enable browsers and pass- credentials against a list of compromised accounts upon each word managers to actively alert users when their specific login, warning the user to change their password if a match is credentials are present in the dataset, protecting their account found. In particular, Google Chrome uses a centralized leak- from the risk of compromise. Indeed, most browsers have age detection service designed by Thomas et al. (USENIX launched some type of password alerting service, automati- Security ’19) that aims to both preserve the user’s privacy and cally checking all credentials entered for prior vulnerabilities. hide the server’s list of compromised credentials. Their inclusion in a browser’s default configuration, how- In this paper, we show that Chrome’s implementation of ever, raises significant privacy concerns from users, as pass- this protocol is vulnerable to several microarchitectural side- words have to be shared with a credential checking service. channel attacks that violate its security properties. Specifically, This prompted Google to incorporate a Private Set Intersec- we demonstrate attacks against Chrome’s use of the memory- tion (PSI) protocol as part of Chrome’s Password Leak De- hard hash function scrypt, its hash-to-elliptic curve function, tection mechanisms [50], removing the need to share users’ and its modular inversion algorithm. While prior work dis- passwords or the server’s list of compromised credentials. cussed the theoretical possibility of side-channel attacks on Another emerging threat to modern systems is the risk of scrypt, we develop new techniques that enable this attack in side-channel attacks. With a plethora of microarchitectural practice, allowing an attacker to recover the user’s password attacks on cryptographic implementations, both from native with a single guess when using a dictionary attack. For modu- code [2, 6, 20, 24, 38, 39, 40, 43, 51, 52, 57], and from the lar inversion, we present a novel cryptanalysis of the Binary browser [3, 26, 28, 29, 37, 48, 53], it is imperative that crypto- Extended Euclidian Algorithm (BEEA) that extracts its in- graphic protocols use side-channel hardened implementations puts given a single, noisy trace, thereby allowing a malicious when processing sensitive information. With Google’s Pass- server to learn information about a client’s password. word Leak Detection protocol [50] using highly customized 1 Introduction cryptographic implementations, in this paper we ask: The past few decades have witnessed a drastic increase in the amount of usernames and passwords leaked via various data Is Chrome’s Password Leak Detection protocol vulnerable to ∗ Work partially done while affiliated with the University of Michigan side-channels? If so, how can attackers exploit the Password † Work partially done while affiliated with the University of Adelaide Leak Detection protocol to recover the users’ passwords? USENIX Association 32nd USENIX Security Symposium 7107 1.1 Our Contribution Section 5 we empirically demonstrate the implications of We analyze Google’s Password Leak Detection protocol, en- side-channel leakage due to rejection sampling by presenting abled by default in Chrome version 106 (latest at the time of attacks on Chrome’s hash to curve implementation, using both writing), and find that Chrome’s Password Leak Detection pro- native and browser-based cache attacks. We also demonstrate tocol leaks information via microarchitectural side-channels. that browser-based cache-attacks are still possible on the latest In particular, by monitoring the cache access patterns while version of Chrome, despite multiple attempts at hardening. running this protocol, we are able to reduce the complexity of Attacking Modular Inversion. As a final contribution, we brute-forcing user credentials such that the attacker often suc- analyze Chrome’s modular inversion operation used during ceeds on the very first login attempt. Since Google’s Password the blinding of the hash digest of the client’s credentials. In Leak Detection protocol is active on default in nearly all mod- Section 6, we attack the Binary Extended Euclidian Algo- ern versions of the Chrome browser, our attack is applicable rithm (BEEA) to show how a malicious password server can to nearly every login attempt on the targeted machine. obtain a digest of the client’s credentials using just a single In our analysis of Chrome’s Password Leak Detection ser- side channel trace. To accomplish this, we developed a novel vice, we found that the client side of the protocol contains cryptanalysis of BEEA that allows an attacker to completely three different components that are vulnerable to side-channel recover the inputs, given only a single, noisy trace. attacks, all leaking independently from one another: (i) using This directly violates the security guarantees of Chrome’s scrypt on users’ input credentials (ii) using hash2curve on the Password Leak Detection, which aims to let clients query the output of scrypt, and (iii) using BEEA to compute the modu- server without leaking information on their passwords. Using lar inverse of the value used to blind the output of hash2curve. our techniques, the server can recover a hash of the client’s The results of these three attacks are summarized in Table 1. credentials, thereby enabling an offline dictionary attack. Attacking Memory Hard Hash Functions. In order to Summary of Contributions: In summary, this work makes prevent attackers from extracting breached credentials from the following contributions. Google’s servers by repeatedly querying the Password Leak • We present the first side channel analysis of Chrome’s Pass- Detection protocol, Google requires that all queries hash the word Leak Detection protocol. checked credentials using a memory hard hash function [44]. • We empirically demonstrate the first end-to-end cache at- This has the effect of slowing down attackers interested in tack against Chrome’s usage of the scrypt, allowing us to extracting the server’s credential list via dictionary attacks, as practically brute force the client’s credentials (Section 4). the computation of the hash digest for each dictionary entry • We present an attack on Chrome’s hash2curve, using both requires a large amount of memory. native code Flush+Reload and browser-based Prime+Probe Unfortunately, Chrome uses scrypt to hash the user’s cre- (Section 5). dentials, a design that is inherently non constant time. By • We present a novel cryptanalysis of the Binary Extended observing scrypt’s input-dependent memory access patterns Euclidian Algorithm (BEEA) that recovers inputs using using Prime+Probe, in Section 4 we demonstrate how attack- only a single noisy trace, and show that its usage in Chrome ers can significantly reduce the cost to brute-force the target’s leaks information about the client’s password, allowing ma- credentials. While prior works have alerted to potential side licious servers to potentially breach the client’s credentials. channel issues with scrypt’s design [11, 25], these approaches (Section 6). fail in practice due to limitations in the bandwidth and accu- racy of the current state-of-the-art side-channel techniques. 1.2 Responsible Disclosure As such, we develop novel techniques that account for noisy We disclosed the vulnerabilities described in this paper to signals with a highly restricted view of the victim’s memory Google through a Crbug report, and shared our paper with a access patterns, resulting in the first end-to-end cache attack number of their engineers. We were able to contact the team against the scrypt algorithm. handling the backend of Chrome’s Password Leak Detection Attacking Hash2Curve. In addition to memory hard service, and suggested and discussed potential mitigations. hashing, Google’s Password Leak Detection protocol also Google stated that they intend to switch to a variant of Ar- requires computing a Private Set Intersection (PSI) between gon2 hash function to mitigate our attack from Section 4, and the client’s credentials and the server’s list of compromised will use a constant-time hash2curve algorithm from [34] to accounts. To that end, Google’s uses a hash to curve algorithm, mitigate our attack from Section 5. At the moment, however, converting the output of scrypt into points on an elliptic curve. Google’s protocol remains unchanged, as Google believes Google then computes the intersection in a homomorphic that given the cost to mount our dictionary attack and the manner, avoiding the need to share user’s passwords or the scope of the threat model, the risk to users is minimal. server’s list of compromised credentials. However, in response to our paper, Chrome mitigated our The hash2curve algorithm, however, uses rejection sam- attack on BEEA from Section 6 through computing the mod- pling [15], which is non-constant time and is explicitly dis- ular inverse by exponentiating by p − 2, thereby removing couraged [35, Appx. A] due to side-channel concerns. In BEEA entirely from Password Leak Detection. 7108 32nd USENIX Security Symposium USENIX Association Section Component Attacker Capabilities Average Entropy Reduction (bits) Required # of Traces 4 Scrypt Native Code 23.41 5 5 Hash-to-Curve Browser Code 0.24 5 6 BEEA Native Code + Blinded Hash Entire Hash1 1 Table 1: Summary of attacker capabilities and results. All three attacks are mutually independent, but can be combined for greater information disclosure. The entropy reduction assumes a dictionary attack from the popular “rockyou.txt" password list, which contains 14,341,564 passwords, or 23.77 bits of entropy. The average is calculated from the experimental outcomes weighted by their probabilities, assuming that the attacker observes the number of login attempts indicated in the final column. CreateDatabase. Before the Password Leak Detection server can handle any requests, it must first hash, blind, and partition its database of leaked credentials. Given the set of credentials, S = {(u0 , p0 ), (u1 , p1 ), ..., (uℓ , pℓ )}, the server first hashes each username:password pair and partitions the database on the n-bit prefix of the hashes of the usernames by computing: S′ = {(H(ui )[0:n] , H(ui , pi )) : (ui , pi ) ∈ S} where H() is the same hash2curve algorithm the client used to compute Q. Then, the server generates its own se- cret key for blinding, b, and blinds the database by computing: S′′ = {(Ui [0:n] , Hib ) : (Ui [0:n] , Hi ) ∈ S′ }. CreateRequest. For each request, the client hashes their username and password pair to an elliptic curve point Q using a hash2curve algorithm, and then blinds Q with a secret key a by computing Qa . The client then constructs request = ′ , Qa ), where U ′ (U[0:n] [0:n] is the prefix of a hash of the username. Figure 1. (Top) Overview of steps in Google’s Password Finally, the client sends request to Google’s Password Leak Leak Detection protocol. (Bottom) User interface displayed Detection server. by Chrome if the verdict is true. CreateResponse. The Password Leak Detection server responds to the request by blinding Qa with b to generate Qab . The server then computes G, the set of blinded credentials 2 Background with username hash prefixes equal to the request’s as: 2.1 Chrome’s Password Leak Detection G = {Hib : (Ui [0:n] , Hib ) ∈ S′′ and Ui [0:n] = U[0:n] ′ }. In order to perform privacy-preserving password checking in the Chrome browser, Google uses a custom protocol called The server then returns the tuple response = (Qab , G). Password Leak Detection [50]. More specifically, Chrome’s Verdict. Upon receiving the response, the client uses Diffie- Password Leak Detection combines anonymity sets, memory- Hellman private set intersection [36] to determine whether hard hashing, and Private Set Intersection (PSI) [36] in order its credentials have been leaked. The client calculates the to check if the given username and password pair is present modular inverse of its secret key a, uses it to unblind the −1 in a data set of compromised credentials, all while preserving doubly blinded hash by computing (Qab )a = Qb , and finally privacy against both malicious clients and malicious servers. checks if Qb ∈ G, indicating compromise. In this section we describe the Password Leak Detection pro- Our focus is on analyzing the side channel leakage from the tocol as implemented in Chrome 106, which differs slighty CreateRequest phase; we now describe this phase in greater from the description given in [50]. detail. See [50] for the complete protocol details. Protocol Overview. Figure 1 presents a high level overview CreateRequest in Detail. Algorithm 1 outlines the Cre- of Google’s Password Leak Detection protocol. The protocol ateRequest phase. The client begins by generating a random consists of four steps, which we now outline: nonce a (Line 2), then creates the string s by canonicalizing USENIX Association 32nd USENIX Security Symposium 7109 1: function C REATE R EQUEST(u, p, n) set of addresses that are “congruent" (i.e. mapped to the same 2: a ← RAND() cache set) with the targeted cache line. 3: u′ ← CANONICALIZE(u) The attacker then primes the cache set by accessing each 4: U ← SHA256(u′ ) address in the eviction set, thereby bringing them into the 5: H ← scrypt(u′ , p) cache. To determine if the victim accessed memory that is 6: Q ← hash2curve(H) 7: Qa ← BLIND(Q, a) congruent with the targeted cache set, the spy process probes 8: U[0:n] ← BYTESUBSTRING(U, n) the cache set by timing accesses to each element in the evic- 9: return (U[0:n] , Qa ) tion set; if any access takes longer than an L3 hit, the attacker infers that the victim brought a cache line into the probed Algorithm 1. CreateRequest: The client uses this function cache set and evicted an element in the eviction set, thereby to hash and blind the username:password pair to send to the revealing the victim’s access to the targeted cache set. server. The deployed version in Chrome uses this algorithm Flush+Reload [57]. This attack has stricter requirements with the prefix length of the username set to n = 26. than Prime+Probe, as it needs the spy and the victim to share memory (e.g., the spy and the victim access a de-duplicated the username u and appending the password p to the canoni- library). In this attack, the spy process prepares the cache by calized username. More specifically, for a username password flushing the targeted cache-line before the victim performs pair (u, p) = (“user@gmail.com”, secret), the canonicalized some action, and then reloads it after the victim finishes. By username and concatenated password is s = “usersecret”. measuring the reload speed, the spy learns whether the victim The client then passes the canonicalized tuple as input brought the targeted cache-line into the cache. Compared to scrypt, a memory-hard hashing algorithm, to compute to Prime+Probe, Flush+Reload samples more quickly, and H ← scrypt(s). This digest is input to Chrome’s hash2curve works with cache-line, rather than cache-set, granularity. algorithm, producing a point Q ← hash2curve(H) on the Flush+Flush [30]. This attack is similar to Flush+Reload, NIST P-256 elliptic curve. Next, the client blinds the hashed with the difference that instead of reloading the targeted cache- point by computing Qa , where we use multiplication to de- line, the spy simply flushes it again and measures the execu- note the elliptic curve group operation. Finally, the client com- tion time of the clflush instruction; a longer time indicates that putes request = (SHA256(u)[0:n] , Qa ) where SHA256(u)[0:n] the cache-line was present in the cache. While Flush+Flush denotes the n least significant bits of the SHA256 hash of u. samples at a rate nearly 7 times faster [30] than Flush+Reload, Comparison With the Original Protocol. Thomas et al. it suffers from a lower accuracy [21]. [50] evaluated a slightly different version of this protocol, 3 Threat Model where the client computes request = (H[0:n] , Qa ), with H[0:n] being the digest resulting from computing scrypt over both In our analysis, we uncovered three separate components the client’s username and password; this has the downside within Chrome’s Password Leak Detection that leave the that it leaks information about the user’s password. To ad- client vulnerable to three separate attacks. Each attack reveals dress this, the authors also propose a Zero-Password Leakage information about the client’s credentials, and the attacks can Variant [50, §3.2] that is similar to Algorithm 1 except that it all be launched independently from one another. As such, uses a memory-hard hash function for computing both U and each attack assumes a different threat model, with differing H (lines 4–5, Alg. 1); the authors cite the cost of a second preconditions, and extracts different amounts of information memory-hard hashing step as a disadvantage. Algorithm 1 from the victim. See Table 1. which is deployed in modern versions of Chrome, does not Attack on scrypt. Following the standard threat model for use a memory-hard hash function to compute U, thus avoiding microarchitectural attacks, in Section 4 we assume that the the downside of the original Zero-Password Leakage Variant. attacker can execute native code under the context of an un- privileged user process on the client’s machine. Furthermore, 2.2 Cache Attacks we assume that the victim is submitting his credentials while A large body of literature examines how two processes can logging into a website on a completely unmodified Chrome inadvertently reveal sensitive information to each other due browser. The attacker then uses side-channels to extract infor- to them sharing the same memory cache [23, 30, 31, 43, 57]. mation on the victim’s execution of scrypt that will reduce the These works show how a victim process’s memory accesses attacker’s search space when conducting a dictionary attack influence the state of the cache, and that an attacking process, on the victim’s input credentials. called the spy process, can deduce what the victim accessed Attack on hash2curve. For our attack on hash2curve (Sec- by indirectly measuring the state of the cache. The three cache tion 5), we again assume that the victim is submitting his attacks relevant to this paper are described below. credentials to a completely unmodified Chrome browser. We Prime+Probe [43]. This attack only requires that the spy relax the assumptions on the attacker, however, and only as- and victim share some level of the cache hierarchy. To carry sume that the attacker has Javascript / Web Assembly code out this attack, the spy first builds an eviction set, which is a running within the victim’s browser. This can occur when 7110 32nd USENIX Security Symposium USENIX Association the victim navigates to a web page controlled by the attacker. 1: function SCRYPT(P, S, N, r, p, dklen) The attacker aims to accomplish the same goal as with the 2: B[0] || B[1] || ... || B[p − 1] ← PBKDF2(P, S, 1, 128 ∗ r ∗ p) scrypt attack: extract information on the victim’s execution of 3: for i = 0 to p − 1 do hash2curve to reduce the search space for a dictionary attack. 4: B[i] = scryptROMix(r, B[i], N) Attack on BEEA. The design of Chrome’s Password Leak 5: DK ← PBKDF2(P, B[0] || B[1] || ... || B[p − 1], 1, dkLen) Detection is such that both client and server are mutually Algorithm 2. Scrypt: The function first uses the PBKDF2 untrusting. That is, even the server should not learn anything key-derivation function to create p blocks each of length 128∗ about the clients’ credentials, and Password Leak Detection r bytes. Each block is then transformed by the scryptROMix is designed with a malicious server in mind. function. The output is the derived key DK. Thus, for our attack on BEEA, we assume that the attacker has access to the blinded output of the hash2curve, which is true when the attacker colludes with the server. We also note means that scrypt is non–constant-time. As noted above, pre- that this access to the blinded point is safeguarded only by vious works [11, 25] have theorized that cache attacks against TLS; as such, any attacker that can compromise the connec- scrypt are possible due to its inherently non–constant-time tion can also access the blinded point. This could occur via nature. Our attack, however, is the first concrete, end-to-end collusion with a TLS middlebox, or even a TLS Enterprise attack on a memory-hard hash function. Root CA certificate installed on the victim’s machine. 4.1 The Scrypt Algorithm For this attack, in Section 6 we assume the attacker has native, unprivileged code running on the victim’s machine, Before describing our attack on scrypt, we now outline the and that he attempts to extract information on the victim’s relevant portions of the scrypt algorithm as it pertains to our execution of BEEA in order to recover the blinding factor. attack against Chrome’s Password Leak Detection. We refer the reader to RFC7914 [45] for a more complete specification. 4 Attacking Scrypt Notation. Following the notation of [45], the scrypt algo- In this section, we explore how to leverage a side-channel rithm takes the following parameters: P, the passphrase to be attack against scrypt as used in Chrome to leak information expanded; S, the salt; N, the CPU and memory cost parameter; about its inputs. We empirically demonstrate how to use a r, the block size parameter; p, the parallelization parameter; combination of cache attacks to recover a small subset of the and dklen, the length of the derived key. accesses into scrypt’s internal memory, which in turn enables Scrypt Overview. Algorithm 2 is an overview of the scrypt an adversary to launch an efficient, offline dictionary attack algorithm. With the exception of P, the password, all parame- against the username:password pairs used as input into scrypt. ters, including the salt, are publicly accessible values that are After examining how scrypt leaks to an ideal side-channel fixed across all users. Thus, the password value P is the only attacker, we relax the requirements and demonstrate how we variable input to scrypt as used in Password Leak Detection. performed the attack in practice, with a much weaker attacker. In Line 2, B, an array of length p where each element is a scrypt in Chrome. Chrome’s Password Leak Detection block 128 ∗ r bytes in length is initialized to the output of the protocol uses scrypt in Line 5 of Algorithm 1, when the client PBKDF2 key derivation algorithm. Then, the loop on Line 3 hashes the username and password together. In this scenario, iterates over each block and replaces it with the value obtained scrypt serves the function of preventing an adversarial client by calling the function scryptROMix(r, B[i], N). Finally, the from efficiently using the Password Leak Detection service password P and the blocks B, along with the desired output to confirm the validity of guessed leaked credentials. length dklen, are passed to PBKDF2 to produce the derived More specifically, the scrypt algorithm [45] is a key- key DK. Most relevant to our attack is Line 4 of Algorithm 2, derivation function (KDF) that is memory-hard [44]. In con- as scrypt’s scryptROMix is highly non constant-time. trast to password hashing algorithms that rely on being com- scryptROMix. The scryptROMix function (Algorithm 3) is putationally expensive, scrypt was designed to require a large responsible for both scrypt’s memory hardness and its non– amount of memory, thus making parallelism impractical, as- constant-time-ness. X is first set to the input block B. Then, the suming it is harder to scale up memory than to scale up com- For-loop at Line 3 initializes V , an array of size N, by setting puting power. This allows Password Leak Detection to better V [i] at each iteration to be equal to scryptBlockMixi (X). We resist attackers employing ASICs or FPGAs to brute force the will refer to this first For-loop as the Initialization Phase. The server’s dataset with massive parallellism. scryptBlockMix function takes an input array of a given size, Achieving Memory Hardness. At a high level, scrypt mixes the bytes, and returns an array of equal size. achieves its memory-hard property by requiring input- The second For-loop, beginning at Line 6, iterates N times dependent random accesses into a very large array. While and makes a non–constant-time, input-dependent memory this forces attackers to store large arrays in memory, it also access (IDMA) each time. We call this second For-loop the 1 This attack recovers all 256 bits of the password hash. In reality, however, Access Phase. The IDMA occurs at Line 8, where j was com- the users’ passwords have far less than 256 bits of entropy. puted by the previous line as j = Integerify(X) mod N. In USENIX Association 32nd USENIX Security Symposium 7111 1: function SCRYPT ROM IX(r, B, N) 2: X ←B 3: for i = 0 to N − 1 do ▷ Initialization Phase 4: V [i] ← X 5: X = scryptBlockMix(X) 6: for i = 0 to N − 1 do ▷ Access Phase 7: j = Integerify(X) mod N 8: T = X ⊕ V [ j] ▷ Input-Dependent Memory Access 9: X = scryptBlockMix(T ) 10: return X Algorithm 3. scryptROMix: The Initialization Phase sets each V [i] to scryptBlockMixi (X). The Access Phase uses j, a function of the input to scrypt, as an index into V , thereby making scrypt non constant-time. Figure 2. While an idealized attacker can observe each mem- ory access indicated by the blue dots, a realistic attack can turn, Integerify(X) simply returns the final 4 bytes of X, inter- only observe memory accesses into a single cache set (yellow preted as a little-endian unsigned integer. Since X comes from boxes) projected onto a single-dimensional trace, represented the output of PBKDF2 via B, which is dependent upon the in- by the red dots. put to scrypt, j is also a function of scrypt’s input. Thus, when j is used as the index into V in Line 8, Algorithm 3 makes an IDMA that is dependent upon the password P, making the when invoking scryptROMix (Algorithm 3). Thus, the leakage entire hashing operation non constant-time. available via a perfect observation of VAP(u, p) is theoreti- cally upper bounded by log2 (40964096 ) = 49152 bits. How- 4.2 Idealized Side Channel Analysis of Scrypt ever, this theoretical limit is further bounded by the size of In this section we perform a side channel analysis of the scrypt the input space into scrypt’s scryptROMix function. Analyz- hashing algorithm assuming an all powerful attacker that can ing the parameter choices for Algorithm 2, we observe that perfectly observe every single memory access into V . Chrome sets p = 1 and r = 8. Thus, the call to the PBKDF2 Memory Layout Mapping. scrypt’s leakage stems from the routine in Line 2 of Algorithm 2 produces a total of 128 · 8 IDMA. As the victim executes the For-loop of the Initializa- bytes of output, mapping passwords to a digest space of size tion Phase, it accesses each memory location in V sequentially 28192 . Finally, given that [50] estimate roughly 23.4–31.2 bil- (Line 4), with a call to scryptBlockMix in between each access lion unique credential pairs, we expect each username and (Line 5). These sequential accesses result in the diagonal line password pair (u, p) to create its own distinct VAP(u, p). of hits comprising the left part of Figure 2. As the elements Credential Recovery via Dictionary Attacks. As each of V are accessed sequentially, an attacker that observes the username and password pair (u, p) creates it own distinct memory accesses in the Initialization Phase can learn which VAP(u, p), an attacker can recover u and p from their elements of V correspond to which memory locations. VAP(u, p) by mounting a dictionary attack. That is, given Obtaining an Input Dependent Access Pattern. The second a plain-text file F of compromised usernames and passwords, half of the accesses in Figure 2 comprises the IDMAs. These an attacker can pre-compute the dictionary accesses are governed by the values of j, which in turn depend DICT (F) := {(u∗ , p∗ ,VAP(u∗ , p∗ )) : ∀(u∗ , p∗ ) ∈ F}. (1) on scrypt’s secret input P, the password. The attacker learned which memory locations correspond to which indexes and Next, during the online phase, in case the attacker ob- can thus correlate these accesses to the Initialization accesses. tains some V-Access-pattern VAP(u, p) corresponding to This allows the recovery of the values of j at each iteration of an attacker-unknown credential (u, p), it is possible to re- the loop in the Access Phase. We refer to the sequence of j cover (u, p) with high probability by performing a search of indexes ( j0 , j1 , . . . , jN−1 ) into V as the V-Access-Pattern. VAP(u, p) in F. Finally, we note that this attack violates the Observing Line 5 of Algorithm 1, the V-Access-Pattern is requirement that the Password Leak Detection server acts as dependent upon the client’s canonicalized user name u′ and an inefficient oracle, as only lookup operations over DICT (F) password p. For a specific user name u and password p we are used during the online phase, while the pre-computation denote by VAP(u, p) the access pattern to V resulting from of DICT (F) from a list of compromised credentials F can be the invocation of Password Leak Detection on u and p. done entirely offline via Equation (1). Leakage Quantification. The precise amount of leakage (in bits) available from VAP(u, p) depends strongly on the 4.3 The Reality of Cache Attacks concrete parameter choices used by Chrome’s Password Leak The previous subsection analyzed scrypt through the lens of Detection protocol. First, we note that Chrome sets N = 4096 a perfect microarchitectural adversary, capable of completely 7112 32nd USENIX Security Symposium USENIX Association reconstructing scrypt’s memory access patterns. While similar approaches were proposed in prior works [11, 25], in this section we outline the challenges in empirically realizing this theoretical attack. As we show, these challenges result in an extremely limited view into the V-Access-pattern for any given scrypt execution, necessitating a different approach. In the following subsections, we outline our solutions and demonstrate the first end-to-end attack on scrypt. Challenge 1: Memory Coverage. The theoretical attack described in Section 4.2 assumes that the attacker can probe Figure 3. On top is what is obtained from averaging across all cache sets in between every iteration of both the For-loops 150 samples of a custom scrypt victim that used the same in Algorithm 3. This would be required to ensure that no memory locations for V each time. In reality, we have to use memory access into V is missed, especially during the IDMA a single-trace, as seen below in red, because Chrome uses where the value of j cannot be predicted by the attacker ahead different memory locations for V every run. of time. In practice however, the loop in the Access Phase executes in less time than it takes to Prime+Probe a set, so A natural approach to mitigating this issue is averaging the an attacker must somehow slow the victim in order to have access time across multiple experiments, as was done in the a chance to probe even a single cache set at each iteration of top graph of Figure 3. However, this is not possible in the the loops. From a signal analysis perspective, this results in case of attacks on unmodified versions of Chrome because the attacker’s view of Figure 2 being limited to accesses to a Chrome’s implementation of scrypt allocates a different set single cache set, i.e. the yellow boxes. of physical memory locations to store V each time. Thus, we Challenge 2: Congruent Cache Lines. There are two yellow developed techniques to overcome the noise and analyze each boxes, indicating that the attacker views multiple indexes experiment in isolation, as opposed to combining the results into V , because the array V is large enough to span multiple across multiple experiments. congruent cache lines. Congruent cache lines are lines that 4.4 Attacking Scrypt in Chrome map to the same cache set, thereby preventing a Prime+Probe attacker from distinguishing between accesses to addresses Having outlined the issues with the theoretical attacks on that map to congruent cache lines. scrypt considered in prior works, we now demonstrate how As Chrome parameterizes scrypt with N = 4096, r = 8, and to attack the scrypt implementation used in an unmodified p = 1, this results in V being an array of 4096 elements with Chrome browser’s Password Leak Detection protocol. each element 1024 bytes in length, spanning 4MiB total. Next, Step 1: Observing Control Flow. In order to observe the as the typical L3 cache on Intel machine uses 0.5 MiB ways, memory access patterns into the V array during every iteration we expect that on average any given element of V will be share of both the For-loops in Algorithm 3, we first need to establish the same cache set with (4096 ∗ 1024)/(0.5 ∗ 1024 ∗ 1024) = a ticker, or a marker that indicates the beginning of a new 8 other elements. Thus, rather than observing a single hit round in either of the For-loops. Since the scryptBlockMix during scrypt’s Initialization Phase followed by a single hit function is called once per iteration, we created our ticker during the Access Phase, the attacker should expect to see by repeatedly using the Flush+Reload attack on a memory 8 hits during the Initialization Phase and an average of 8 addresses holding the code of this function. We will use this corresponding hits during the Access Phase. ticker as an indicator of start every iteration of the For-loops Empirically demonstrating this issue, we executed an in- in Algorithm 3, allowing us to assign every subsequent side strumented version of Chrome’s scrypt code while monitoring channel observation to its corresponding loop iteration. the accesses to a single cache set. The top graph in Figure 3 Step 2: Performance Degradation. Next, as the iterations illustrates the access times, where we added the results of of the For-loops in Algorithm 3 are so short in duration, we 150 scrypt traces corresponding to the same cache set. The y- also had to slow the execution of the scryptBlockMix function axis measures the number of times a cache-hit was observed, using a performance degradation attack [8]. To that aim, we while the x-axis indicates which round it was recorded in. In repeatedly used the clflush instruction in order to flush a code this example, there are roughly 9 peaks for the Initialization region corresponding to the Salsa 20 Core function, which is Phase followed by another 9 for the Access Phase, which is repeatedly called inside scryptBlockMix. close to what is expected. Step 3: Prime+Probe. With the ticker and performance- Challenge 3: Noise. Further complicating the attack is the degradation established, we choose a random cache set and presence of noise, which prevents us from perfectly learning mount a Prime+Probe attack on it after each occurrence of which n indices are accessed at which rounds in the Access the ticker. As outlined in Section 4.3, the memory layout of V Phase. This can be seen in the Figure 3(bottom), which shows inside the CPU’s cache implies that mounting a Prime+Probe a single trace containing 80 hits, though only 16 are expected. attack on a given cache set discloses when an access is per- USENIX Association 32nd USENIX Security Symposium 7113 formed to any of the roughly 8 congruent elements of V , with- we empirically found it optimal to expanding our scoring al- out the ability to further distinguish between the elements. gorithm to also consider the indexes that match in the low 2 bits immediately above and below where the hit occurred. 4.5 Handling Noise Determining the Transition Point. Next, after denoising the As outlined in Section 4.3, a further issue that complicates our ticker during the Initialization Phase of Algorithm 3, we must attack is the presence of noise, which takes the the form of off- locate the “halfway point": the point in Line 6 in Algorithm 3, sets in the ticker and additions or deletions in the Prime+Probe just before the second For-loop starts. This allows us to realign cache hits. As noted above, we cannot simply average out the trace after drifting for 4096 rounds in the Initialization the noise over multiple traces, since Chrome’s scrypt imple- Phase. We do so by exploiting the fact that one iteration of mentation ends up using a different set of physical memory the loop in the Initialization Phase (Line 3) executes more locations to store V each time. Thus, rather than combining quickly than one iteration in the Access Phase (Line 6) due several measurements into a clean trace, we instead overcome to the additional code at Lines 7 and 8. Thus, by looking the noise while analyzing each trace in isolation. for a point where the time between ticker hits consistently Prime+Probe Noise. When we conduct the Prime+Probe increases, we are able to identify the halfway point, allowing attack on a cache set, we see a large amount of false positive us to identify the ticker’s transition to the Access Phase. noise, where we record cache hits during rounds where there Denoising the Ticker During Accesses. Ticker noise during should not be. In addition, we observed a minimal amount of the Access Phase has a more straightforward solution. When false negative noise, where cache hits are missing. searching for the hits in the Access Phase that correspond to To overcome this, we implement a scoring system for our the indices found during the Initialization Phase, the attacker dictionary attack, where we assign points to candidate pass- simply expands her search for any hit within 10 rounds of words based off how well they “fit" the results from a trace. the expected location. We empirically found 10 to be a good Given an index j0 that is accessed during the Initialization compromise between being too small to overcome the noise, phase, the attacker can pre-compute at which rounds during and being so large as to generate too many false positives. the Access Phase V [ j] will be accessed in case the password Avoiding Averaging. Overall, we were able to combine candidate is correct. For each of those rounds in the Access the above denoising and scoring techniques into an algorithm Phase which the Prime+Probe trace contains a memory ac- to identify candidate passwords based on how well they fit cess, the candidate password’s score is incremented by 1. the limited, noisy information the attacker gained on the V- We repeat the above approach for each index, Access-Pattern. This improves on prior theoretical attacks on j0 , j1 , ..., jn−1 , for which accesses are detected during scrypt [11, 25], which assumed the attacker has a noiseless Initialization Phase, in order to compute the candidate’s final and perfect access to the V-Access-Pattern. score. Finally, after applying this approach on all password candidates in the dictionary, we rank the highest scoring 4.6 Empirical Evaluation candidates as the most likely passwords. We now evaluate the effectiveness of our attack on Chrome’s Ticker Noise. While the above approach is useful for han- Password Leak Detection protocol. dling noise present in the accesses to V obtained using the Experimental setup. We conducted our attack against an Prime+Probe channel, we must adapt this algorithm to also unmodified Chrome binary running on an Acer Aspire E 15 account for noise that is present in our Flush+Reload ticker. laptop, equipped with 8GiB of DDR4 memory and an Intel We begin by recalling that we use the ticker to determine i5-8250U CPU. The i5-8250U features 4 cores and 8 threads at which round the accesses into V are made, during both and is equipped with a 6MiB 12-way set associative L3 cache. Initialization and Access phases. Due to both false negatives Our machine was running Ubuntu 20.04.3 with Linux kernel and false positives, the round corresponding to the memory version 5.8.0-44. access is unfortunately rarely correct. Attack Scenario. Assuming an attacker with native un- Denoising the Ticker During Initialization. We observe privileged code execution on the target machine, we imple- that ticker noise is most damaging during the Initialization mented the attack described in this section using the Mastik Phase. If the ticker is off, then our algorithm ends up assign- toolkit [56]. We ran our attacker against an unmodified ver- ing points for the trace fitting the wrong V-Access-Pattern. sion of the Chrome browser, and evaluated our attack against Fortunately for the attacker, however, the accesses to V during 10 randomly chosen passwords from the “rockyou.txt" dic- the Initialization Phase occur deterministically, with elements tionary. For each password, the victim browser submitted the of V accessed sequentially (see Line 4 of Algorithm 3). Thus, username:password pair of (“z", pw) into a website 5 times, the attacker learns some of the low bits of the index due to where pw was the randomly chosen password from “rock- the elements of V being smaller than one page. In particu- you.txt". This resulted in our side-channel attacker obtaining lar, each element of V is 1024 bytes, meaning that 4 fit into 5 traces for each password the victim submitted, where each each page, and for each ji accessed during the Initialization trace is the result of using side-channels to leak information Phase, the attacker learns ji (mod 4). With this optimization, about the victim’s V-access-pattern. 7114 32nd USENIX Security Symposium USENIX Association A Dictionary Attack. With the collected traces in hand, we 1: function hash2curve(m) applied the approach described in Section 4.5 and conducted 2: px ← RandomOracleSHA256(m) an offline dictionary attack. We follow the precedent of pre- 3: while !OnCurve(px ) do ▷ Input-Dependent Loop vious works and benchmark our attack on the “rockyou.txt" 4: px ← RandomOracleSHA256(px ) password file, which contains 14,341,564 plaintext passwords 5: return (px , py ) stolen during the 2009 RockYou data breach. Algorithm 4. Google’s hash2curve implementation. We computed the V-access-pattern for every entry in the entire “rockyou.txt" file and scored them against the 10 sets of traces corresponding to the 10 passwords the victim submit- elliptic curve. Such a primitive is useful for a number of ted. We ran this computation on an Intel Xeon server machine, cryptographic applications; in the case of Password Leak featuring a Platinum 8352Y CPU, 128 cores, and 1.8TB of Detection, Chrome’s hash2curve algorithm serves the purpose memory. It took about 8400 core hours of offline computa- of mapping the output of scrypt to a point on an elliptic curve tion to complete the dictionary attack, or about 3 days when to prepare it for use with Diffie-Hellman PSI [36]. This can parallelized across all 128 cores. be seen in Line 6 of Algorithm 1. The results of our end-to-end attack are displayed in Fig- Chrome’s Hash2Curve Implementation. We describe the ure 4. The correct password ended up scoring higher than hash2curve algorithm used in Password Leak Detection in all other passwords in the dictionary 80% of the time. Thus, Algorithm 4. The variable px is first computed by passing if the attacker were to attempt to log in with the candidate the string m as input to the function RandomOracleSHA256, passwords in descending order, the attacker would success- which repeatedly uses SHA-256 and modular addition to fully log into the victim’s account on the very first try, 80% approximate a pseudo-random function (PRF). The output of the time. In the worst case, the attacker would succeed on px is then repeatedly updated by being assigned the value the 6th try. Finally, we note that with such a low number of RandomOracleSHA256(px ) until px is a valid x coordinate attempts required for success, no reasonable amount of rate of a point P = (px , py ) on the NIST P-256 curve (Line 3). The limiting on password attempts can prevent our attacker from algorithm then outputs P = (px , py ) (Line 5). About half of compromising the target’s account. the possible values of x correspond to points on the curve, so the while-loop in Line 3 of Algorithm 4 terminates with 8 probability 1/2 in each iteration. Side Channel Vulnerability of Hash2Curve. While Frequency 6 4 constant-time hash2curve algorithms do exist [54], we note 2 that the hash2curve algorithm used by Chrome’s Password Leak Detection is inherently not constant-time. More specifi- 0 1 2 3 4 5 6 cally, Algorithm 4 uses a rejection sampling method, repeat- edly iterating over candidate px values until a suitable value Number of login attempts until success is found. While this design pattern was originally proposed by [15], such an implementation is explicitly discouraged by Figure 4. Histogram of the results from our attack on scrypt in [35, Appendix A] due to side channel considerations. an unmodified Chrome browser. 80% of the time, the attacker A Dictionary Attack on Hash2Curve. An attacker can guesses the correct password on the very first attempt. exploit the rejection sampling design of Algorithm 4 for mounting dictionary attacks similar to those mounted in Sec- tion 4. Given a credential dataset D, an attacker can apply the 5 Attacking Hash2Curve hash2curve algorithm to every entry of D, obtaining the corre- Moving away from attack Password Leak Detection scrypt im- sponding number of iterations taken by Line 3 of Algorithm 4, plementation, in this Section we will examine how Chrome’s since it is an input-dependent loop (hereinafter IDL). Next, by hash2curve usage reveals bits of the user’s credentials. Before using a side-channel attack to obtain the number of iterations demonstrating end-to-end attacks from both native code and taken by the IDL on the target’s credentials, the attacker can the Chrome browser on this part of the Password Leak Detec- eliminate candidates of D that do not induce the same number tion protocol, we now proceed to review Google’s hash2curve of iterations, thereby reducing the attack’s search space. construction and implementation. We used the unmodified Chrome version 106, the latest at the time of writing, for all 5.2 Native Attack on Hash2Curve analyses and experiments in this section. In this section we describe a Flush+Reload based attack on Chrome’s hash2curve algorithm, executed from unprivileged 5.1 Hash2Curve Overview native code running in the target’s machine. Here, we empir- The hash2curve algorithm takes an arbitrary length input ically found the strongest results when using Flush+Reload string and deterministically outputs a point on a specified on a string that is touched by Chrome’s GetPointByHashing- USENIX Association 32nd USENIX Security Symposium 7115 5.3 Attacking Hash2Curve Within Chrome Having established the vulnerability of Chrome’s hash2curve algorithm to Flush+Reload attacks from native code, in this section we present a browser-based attack on Chrome’s hash2curve implementation. More specifically, we attack hash2curve from within an unmodified Chrome browser, us- ing an attacker webpage that executes malicious JavaScript and WebAssembly code. Figure 5. Flush+Reload attack accuracy as a function of the This is a weaker assumption on the attacker’s capabiltiies number of hash2curve loop iterations. than in Section 5.2, as it is usually easier for an attacker to trick a victim into visiting the attacker’s web page. As such, modern browser versions recognize the danger of browser- ToCurveInternal function, which corresponds to the OnCurve based side-channels, and employ a heavily sandboxed envi- test within the IDL. ronment for code execution compared to a native scenario. Attack Outline. To mount a Flush+Reload attack, we exe- Accordingly, browser-based adversaries face additional tech- cute an unprivileged attacker process monitoring Chrome’s nical challenges, which we now describe. GetPointByHashingToCurveInternal function on the target ma- Flushing Data in the Cache. The Flush+Reload tech- chine. We then open an exemplary login page and complete nique used in Section 5.2 requires the clflush instruction and the login process with a set of user credentials. This triggers shared memory between attacker and target. However, neither Chrome’s Password Leak Detection protocol, allowing us to is available in a browser environment. monitor the exact number of iterations made by the IDL. Instead, we use Prime+Probe to observe the activity of Attack Evaluation. In order to measure the accuracy of our Chrome’s hash2curve algorithm, using the work of Vila et Flush+Reload attack, we gathered 5500 pairs of username and al. [53] in order to efficiently generate eviction sets. password, where every 500 pairs generate the same number High-Precision Timing Source. Measuring the cache access of iterations of the IDL between 0 and 10 (inclusive). We then patterns of the hash2curve algorithm requires the attacker to executed the attack on each credential pair, noting the amount distinguish cache hits from misses, necessitating a timer with of detected iterations compared to the ground truth. a precision of tens of nanoseconds. However, modern Chrome Analyzing the results, our Flush+Reload attack was able versions deliberately limit the timer resolution to 5 µs, aiming to correctly identify the number of loop iterations for 4960 to foil side channel attacks [55]. credentials, resulting in a total accuracy rate of 90.18%. Next, We sidestep this issue by using the SharedArrayBuffer in Figure 5 we outline the accuracy of our attack as a function JavaScript API, which provides a primitive for a precise count- of the actual loop iterations performed by the IDL. ing thread on the order of nanoseconds [47]. While SharedAr- Finally, as the IDL exists with probability 1/2 for each loop rayBuffer was previously disabled by Chrome in an attempt iteration, we can use the data depicted in Figure 5 to compute to mitigate speculative execution attacks, recent versions of our weighted success probability as Chrome re-enabled this primitive due to the presence of dedi- cated Spectre countermeasures [12, 46]. W[SR] = 1 ∗ 1/2 + 0.954 ∗ 1/4 + · · · + 0.826 ∗ 1/211 Just-In-Time Compilation. In contrast to attacks that are mounted using native code, which have near-complete control which is roughly equal to 97%. over the attack code executed by the target machine, browser- A Dictionary Attack. By recovering the number of based adversaries are limited to code emitted by Chrome’s iterations with 97% accuracy, the attacker is able to conduct JavaScript and WebAssembly execution engines [13, 49]. a dictionary attack and filter out candidate passwords that This introduces measurement noise, making traces obtained don’t have the same number of iterations. As the IDL exists through side-channels unreliable and nondeterministic. with probability 1/2 for each loop iteration, the attacker To overcome this issue, we observe that it is possible essentially learns an additional bit about the password with to introduce a warmup stage into our attack code, causing each iteration; this means that in the ideal case where the Chrome to always run its optimizing compiler over our high- attacker can perfectly recover the number of hash2curve level Prime+Probe implementation. Furthermore, we initial- loop iterations, the expected number of bits by which the ize our code in a way that Chrome’s optimizing compiler password’s entropy is reduced is: will cache its output [17], allowing us to consistently use Prime+Probe across many attack runs. With both measures in ∞ 1 1 1 place, we achieve a greater probing frequency compared to E[B] = 1 ∗ + 2 ∗ + 3 ∗ = ∑ i/2i = 2. 2 4 8 i=1 naive Prime+Probe implementations, allowing us to reliably monitor the execution of the IDL. 7116 32nd USENIX Security Symposium USENIX Association Automatically Selecting the Correct Eviction Set. Having generated eviction sets using Vila et al. [53], we must now determine the eviction set corresponding to the IDL. To that aim, our attacker page renders an attacker-controlled login page inside an iframe, populating it with dummy credentials known to the attacker. This triggers the execution of Chrome’s Password Leak Detection protocol, eventually resulting in invocations of the IDL. For each eviction set generated by Vila et al. [53], we exe- cute a Prime+Probe attack on our dummy credentials using the above procedure, locating the eviction set that recorded the Figure 6. Elapsed counting thread ticks over Prime+Probe correct amount of cache misses corresponding to the execu- iterations, with the 15 spikes corresponding to the 15 IDL tion of the IDL. With the correct eviction set in hand, we can iterations highlighted in red. now mount Prime+Probe attacks on the target’s credentials. Attack Evaluation. Our goal is to determine whether browser-based attacks on Chrome’s hash2curve algorithm are capable of mirroring the performance of native attacks outlined in Section 5.2. With this in mind, we mount a Prime+Probe attack on a targeted username and password pair, using the eviction set found earlier. Figure 6 presents a time series of Prime+Probe attack iter- ations; the y-axis plots the number of counting thread ticks elapsed while accessing the eviction set in each iteration. We note the 15 spikes of probes that have at least 500 ticks (in red), Figure 7. Effect of increasing the login attempts on the num- which corresponds to 15 iterations of the IDL when Chrome’s ber of successes for each number of IDL iterations. Password Leak Detection is invoked on our credential. We can distinguish the spikes resulting from the target string be- ing accessed from system noise, since the latter only results 6 Attacking Blinded Hashes in access times no longer than 200 ticks. Finally, Figure 7 shows how many times out of 30 we were able to observe We now examine how Chrome blinds the client’s credentials the correct number of spikes when running our attack during before sending them to the server. In particular, we found 1, 5, and 10 login attempts. While we could not distinguish that Chrome uses the Binary Extended Euclidian Algorithm targeted credentials inducing less than 5 IDL iterations, we (BEEA) to compute the modular inverse of its secret key for observe that 5 logins suffice for the attack to succeed in at blinding, which is susceptible to side-channel attacks. least half the trials. This results in our browser-based attack To investigate the extent to which this compromises Pass- with 5 traces reducing the password’s entropy by an expected: word Leak Detection’s security guarantees, we developed a 19 19 20 E[B] = 5 ∗ 30∗2 5 + 6 ∗ 30∗26 + ... + 10 ∗ 30∗210 = 0.24 bits. novel cryptanalysis technique that can recover BEEA’s in- While this may seem like a trivially small amount of leak- puts after observing only a single, noisy trace obtained via a age in the average case, we contend that only examining cache side-channel. We then used our technique to success- the average case and neglecting to account for the danger to fully demonstrate the first ever microarchitectural single-trace passwords with higher numbers of IDL iterations belies the attack on the BEEA algorithm, improving on prior work that severity of our attack. Instead, we emphasize that our results required a controlled execution environment (e.g. SGX [41]) concretely demonstrate the risk of a significant amount of or constraints only present during RSA key generation [7]. leakage for a non-trivial number of cases; as shown in Fig- Blinding Chrome’s Password Leak Detection Protocol. ure 7, when there were 10 IDL iterations and 5 login attempts In the Password Leak Detection algorithm, after the client were observed, the attack succeeded in 20 of 30 trials. This hashes the user’s credentials to the point on the curve Q, the means that for these passwords, the attack reduced the entropy client must blind the hash before including it in the request of the password by 10 bits, 2 out of 3 times. to the server. This takes place on Line 7 of Algorithm 1. Since 1 out of 1024 credentials result in 10 IDL iterations, Specifically, the client blinds the hash by computing Qa on and the attack succeeded 2 out 3 times, this means 1 out of the elliptic curve, where a is the secret key. This blinding 1536 passwords will leak 10 bits. Given the scale at which a serves the purpose of concealing all bits of the client’s hashed browser-based attack can be launched, that fact that our attack credentials from the Password Leak Detection server. can leak 10 bits from 1 in 1536 passwords is concerning, even This step is in fact critical, as the Password Leak Detection if the average password leaks very little. protocol was designed to preserve privacy in both directions, USENIX Association 32nd USENIX Security Symposium 7117 meaning that the server should learn nothing about the client’s 1: function BN _ MOD _ INVERSE _ ODD(a, n) credentials. If the server learns the hash of the client’s creden- 2: A ← n, B ← a, X ← 1, Y ← 0 tials, the server can then launch a dictionary attack to brute 3: while B ̸= 0 do force the client’s plaintext password. 4: while even(B) do 5: B ← B/2 Computing Modular Inverses using BEEA. To unblind 6: if odd(X) then ▷ Branch 1 the response received from the server, the client also needs to 7: X ← X +n ▷ Branch 5 compute a−1 mod p, where p is the prime modulus used for 8: X ← X/2 the NIST P-256 elliptic curve. The value of a−1 mod p is then 9: while even(A) do used to unblind the response received from the server in order 10: A ← A/2 to complete the Diffie-Hellman PSI [36]. Next, to perform 11: if odd(Y ) then ▷ Branch 2 the computation of a−1 mod p, Chrome uses the BEEA algo- 12: Y ←Y +n ▷ Branch 5 rithm, as implemented by BoringSSL’s BN_mod_inverse func- 13: Y ← Y /2 tion, which in turn calls BoringSSL’s BN_mod_inverse_odd 14: if B ≥ A then to take advantage of an optimization for odd moduli. 15: X ← X +Y ▷ Branch 3 Threat Model. Similarly to Section 4’s threat model, we 16: B ← B−A also assume a side-channel adversary that is able to run un- 17: else privileged native code on the victim’s machine. Moreover, 18: Y ←Y +X ▷ Branch 4 in this section we assume that the attacker has access to the 19: A ← A−B blinded hash, and wants to obtain the value of the unblinded 20: return Y hash. Access to the blinded hash could occur in practice if the Algorithm 5. Binary Extended Euclidian Algorithm: Pseu- server participating in the Password Leak Detection protocol docode for Chrome’s BEEA implementation, which is opti- colludes with the attacker; unblinding the hash would allow mized for the odd modulus n used in NIST P-256. We make the server to brute force the client’s credentials via a dictio- the observation that the conditional add, labeled as Branch 5, nary attack, a scenario Chrome’s Password Leak Detection allows for error correction in a noisy trace. protocol was specifically designed to protect against. We note that the attacker can access the blinded hash section we develop a novel noise-tolerant single-trace attack in other ways, besides colluding with the server; if the at- on the BEEA algorithm that enables extraction of modular tacker could somehow compromise the TLS connection (e.g. inverses computed by BEEA. through a TLS Enterprise Root CA certificate, or a TLS mid- dlebox), the attacker would gain access to the blinded hash. 6.2 Attacking BEEA 6.1 Prior Attacks on the BEEA Algorithm At a high level, our attack uses cache-attacks to determine the The BEEA Algorithm has been studied extensively from control flow of the BEEA, which in turn allows for the recov- a side-channel perspective due to the widespread need for ery of the inputs to BEEA. In this case, the inputs to BEEA computing modular inversions in cryptographic systems (e.g. are a, the client’s blinding key, and n, the prime modulus for RSA, ECDSA). While prior side-channel attacks against the elliptic curve NIST P-256. BEEA are well known, the specific manner in which Google Chrome’s BEEA Algorithm. Algorithm 5 is a pseudocode uses BEEA in Chrome’s Password Leak Detection proto- of the BEEA algorithm. Borrowing notation from [1], we col prevents the application of prior attack techniques. More use SHIFTS[i] to denote how many times the branch at Line specifically, the theoretical analysis of BEEA done in [1] as- 4 or Line 9 was taken at the ith iteration of the outer while- sumes that the attacker can obtain perfect, noiseless traces loop at Line 3 (only one branch or the other will be taken of the BEEA’s execution, which is not possible with current during any given iteration, as one of A and B will be even attack techniques. Furthermore, Chrome generates a new ran- and the other odd at the beginning of each iteration). We dom blinding factor a upon each generation of a request, let SUBS[i] denote the outcome of the comparison at Line thereby precluding combining information across traces via 14, such that SUBS[i] = 3 if branch 3 is taken, and 4 if if either averaging or lattice attacks [5, 27]. branch 4 is taken at the ith iteration. We note that for any Thus, in order to attack BEEA as it is used in Chrome, a iteration i, if SUBS[i − 1] = 3, then the SHIFTS[i] must all side-channel attacker must operate with only a single, noisy take place in Branch 1 due to the subtraction at Line 16 trace of the BEEA execution. In prior works, this has only in Algorithm 5. Likewise, if SUBS[i − 1] = 4, then the next ever been accomplished by either placing the victim inside iteration’s SHIFT S[i] take place in Branch 2. of an SGX enclave [41] or by exploiting the redundancy and Perfect Trace Requirement. Acıiçmez et al. [1] previ- relations between various known parameters when BEEA ously showed that if an attacker can perfectly recover the is used during RSA key generation [7]. As neither of these SUBS[] and SHIFTS[] for all iterations, they can reconstruct scenarios apply to Chrome’s Password Leak Detection, in this both inputs to BEEA in polynomial time. The downfall of 7118 32nd USENIX Security Symposium USENIX Association 1: function B RANCH AND P RUNE(Array Trace) Probing the BEEA Algorithm. While prior works [1, 7, 2: pq ← PriorityQueue() 27, 41] only probed Branches 1, 2, 3, and 4 via cache attacks, 3: pq.push(100, 0, Key([])) we make the observation that detecting the conditional add in 4: while notEmpty(pq) do Branch 5 can be used to prune potential keys during our search 5: (score, i, curKey) ← pq.pop() algorithm. This is because the values X and Y are completely 6: if i == len(T ) then determined by all the previous branches taken. Furthermore, 7: Out put(curKey) 8: Continue since X and Y start off initialized to 1 and 0 respectively, and 9: curBranch ← Trace[i] the value n is known to be the prime modulus of NIST P-256, 10: newKey ← curKey.append(curBranch) then if we know all the prior branches, we can determine if 11: if correctXY (newKey) then branch 5 can possibly be taken at the current iteration. We 12: if curBranch == 5 then will now cover in detail exactly how our branch-and-prune 13: score ← score + 20 algorithm works in Algorithm 6. 14: pq.push(score, i + 1, newKey) Algorithm Description. In Line 1, our algorithm takes 15: if LastSub(curKey) == 3 then Trace, an array of branches corresponding to the sequence of 16: newKey ← curKey.append(1) branches taken by BEEA, as input obtained through a side- 17: X ← CalculateX(newKey) channel attack. The branches are simply numbers from 1 18: if isOdd(X) then through 5, corresponding to the labeled branches in Algo- 19: Continue rithm 5. We make the assumption that Trace only contains 20: pq.push(score − 20, i, newKey) deletion errors, and only deletions of 1s and 2s occur; further- 21: else 22: newKey ← curKey.append(2) more, 1s and 2s are never deleted when they are immediately 23: Y ← CalculateY (newKey) preceding a 5. In Section 6.4 we explain why this was the case 24: if isOdd(Y ) then for when we used cache-attacks to obtain traces on BEEA. 25: Continue The data structure that we use to process our candidate keys 26: pq.push(score − 20, i, newKey) is the priority queue on Line 2, which is sorted by the score of Algorithm 6. Branch and Prune: Pseudo-code for our each candidate key. A key’s score is a measure of how close branch and prune algorithm that recovers the complete in- to the real key we believe the candidate key to be. In Line 3, put to BEEA, given a single noisy trace. we populate pq with with a key with score equal to 100, an iterator i equal to zero, and an array of branches equal to 0. The iterator is used to track how far along the trace the key has this analysis of BEEA is that a single error in either SUBS[] progressed through. The key’s array of branches represents the or SHIFTS[] completely foils the recovery of the inputs. More- sequence of branches taken by the BEEA algorithm. As the over, the attacker cannot determine if there were any errors, algorithm progresses, we will incrementally build up longer and thus whether or not the result is correct. keys that get progressively closer to the real key. Prior work [41] that utilized the analysis of [1] was able to Finding New Keys. At each iteration of the loop on Line extract BEEA’s inputs via perfect side channel traces, care- 4, we process the highest scoring key, corresponding to the fully controlling its execution within an SGX enclave. For at- key that we currently believe to be the closest to the true key, tacking Chrome, however, the traces obtained via Flush+Flush and push additional keys onto the priority queue with one contain a substantial number of errors over the course of additional branch added at a time. BEEA’s roughly 700 branches on its randomized inputs, pre- To generate these additional keys, our algorithm branches venting us from applying the analysis of [1]. Thus, we over- at each iteration of the outer while loop to create 2 additional come the perfect trace requirement by presenting the first keys. The first new key is formed by branching towards the analysis of BEEA input recovery for the case of noisy traces. Trace by appending the next branch within Trace to the cur- rent key, as seen on Lines 9 and 10. This candidate key is then 6.3 Cryptanalysis of a Noisy Trace pruned on Line 11 if it fails to satisfy correctXY (). To correct the errors present in a noisy BEEA trace, we draw Pruning on X and Y . The correctXY () function inspects inspiration from prior works on partial key recovery [32, 33] the newKey’s array of branches. Assuming an execution of and develop a branch-and-prune style algorithm for BEEA. BEEA that follows those branches, correctXY () then deter- At a high level, this involves searching for the correct key, mines if it is possible for all occurrences of Branch 5 to be at where a key is a sequence of branches that could have been the locations that they are. That is, after every occurence of taken by the execution of BEEA. Our algorithm repeatedly Branch 1 and Branch 2, it checks to see that X or Y , respec- branches towards the most probable keys, as determined by tively, at that point are odd if and only if a Branch 5 occurs on how closely they align with the trace. We then exploit the the subsequent branch. If not, then correctXY () returns false, relationship between different segments of the keys to prune and the key is pruned. On the other hand, if the key is not key candidates, until the correct key is found. pruned, then newKey’s score is incremented by 20 on Line USENIX Association 32nd USENIX Security Symposium 7119 13, its position in Trace is incremented by 1, and the key is 6.4 Implementing the Attack pushed onto pq on Line 14. In this section we describe how we implemented our cache We note that in actuality, only the most recent branch and attack to obtain a trace against BEEA, and how our branch- the resulting X and Y need to be checked in this manner. and-prune algorithm recovered its inputs. Any inconsistencies between X and Y and the sequence of Software Setup. Chrome is statically linked against Bor- branches earlier in the key would have resulted in that key ingSSL, and as such the Password Leak Detection logic already having been pruned. calls BoringSSL’s BN_mod_inverse function to compute the modular inverse of the blinding factor. To benchmark False Positives. While it is possible for correctXY () to our attack, we developed a test harness that calls Bor- return true for keys that do in fact have errors in them, as ingSSL’s BN_mod_inverse, compiled with gcc version 9.4.0 the distance from the errors grows, the chance of continual using an -O0 flag. Mirroring Chrome, we call BoringSSL’s false positives is equal to 1/2n , where n is the number of 1 BN_mod_inverse(a, n) with random 256-bit values of a, where branches and 2 branches since the error. This is because X n set as NIST P-256’s prime modulus. and Y are modified with each occurrence of branches 1 and This is in contrast to the prior two attacks on scrypt and 2, effectively randomizing whether or not they are odd at any hash2curve, where we conducted our attacks against unmod- given point. Furthermore, the subtractions in branches 3 and ified versions of Chrome. We did this because BEEA sees 4 ensure that errors in X propagate to errors in Y , and vice widespread deployment across numerous commonly imple- versa. As n grows, the probability of continuing to follow an mented cryptosystems, and we believe that our novel crypt- incorrect key becomes vanishingly small, and the incorrect analysis has implications on these as well. Thus, there are key values are pruned back to where the error occurred. broader impacts in analyzing how BEEA leaks using our Inserting Branches. After branching towards the Trace, attack technique. By benchmarking our attack against Bor- the second additional key is found by inserting a potential ingSSL directly, we demonstrate that our attack applies to branch into curKey. This is how keys that contain the branches a wider variety of BEEA usages (any binary that uses Bor- deleted by the noisy trace are discovered and added to pq. ingSSL’s implementation, such as Chrome), and does not rely Since the Trace obtained in Section 6.4 only contains dele- on nuances specific to Chrome. tions of 1 Branches and 2 Branches, the LastSub() function at Finally, we run experiments on a laptop featuring a Quad Line 15 only needs to determine whether to insert a 1 Branch Core Intel i5-8250U CPU, and 4 GB of RAM. or a 2 Branch, which depends on whether the most recent Flush+Flush Probing Locations. In order to generate a trace branch is 3 or 4 respectively. that attempts to reconstruct the control flow of the the victim’s After inserting the potential branch at Line 16, Line 17 uses BEEA execution, we used the Flush+Flush attack [30] in order the function CalculateX() to determine the value of X within to monitor the 5 branches marked in Algorithm 5. We note BEEA after executing the branches in newKey. If this X is odd, that doing so requires a total of 4 Flush+Flush probes, as the this would induce a Branch 5 to follow. However, since there same probe can be used to monitor both Branch 1 and Branch are no deletions of Branch 5 or the immediately preceding 2 since these share the same call to BN_rshift1 on Lines 8 and branch 1s and 2s in Trace, inserting an additional branch 5 13. Similarly, a single probe monitors Branch 5 on both Line automatically renders the key incorrect; as such, we prune the 7 and Line 12 as they make the same call to BN_uadd. We use key at Line 19 if X is odd. Otherwise, we push the newKey the remaining two probes to detect branches 3 and 4, which onto the priority queue at Line 20. To prioritize keys that in turn allows us to discern between branch 1 and branch 2, align more closely with Trace, we decrement the newKey’s as described in Section 6.2. score by 20, since for most branches the Trace is correct and Signal Amplification via Core Assignment. As our does not require an insertion. We leave i untouched because i5-8250U processor has 4 physical cores, we run each we only inserted an additional branch, and did not progress Flush+Flush probe on a separate core, while having the probe through Trace. Lines 22 through 26 serve the same purpose, for Branch 5 running on the sibling virtual core to the process only for when LastSub == 4. executing the BEEA algorithm. As the BEEA process now shares its caches with the probe for branch 5, due to [4] this re- Termination. Once i iterates through the entirety of the sults in the signal for branch 5 becoming unmistakably strong. Trace, the candidate key is output, along with its score. The This is important, as it virtually eliminates false positives or algorithm can continue to run indefinitely, continuously out- negatives for branch 5, allowing us to prune keys aggressively putting more complete keys as it explores them. The higher by error correcting with the occurrences of Branch 5. Further- the score of a key, the more likely it is to be the correct one. more, this also results in a trace where branch 1s and branch Intuitively, the true key is the one that aligns most closely 2s immediately preceding branch 5s are not deleted in the with Trace, with the insertions in the correct places that result trace, as branch 5 can only ever take place after a 1 or 2. in X and Y being odd whenever dictated by the occurrences Gathering Traces. We then allow the BN_mod_inverse func- of branch 5 in the Trace. tion to run while the probing processes monitor the branches. 7120 32nd USENIX Security Symposium USENIX Association We parse the resulting data from the Flush+Flush probes form Scrypt. Chrome uses scrypt as its hash algorithm for Pass- the trace, which is the sequence of branches within BEEA ob- word Leak Detection due to its memory-hardness properties. served by the attacker. We find that this results in no insertions A memory-hard hash function is one where the cost of eval- or deletions of branches 3, 4, and 5. However, it is common uating the hash function is primarily dominated by the cost for there to be insertions or deletions in the number of branch of memory, as opposed to the cost of compute power. While 1s and branch 2s in each round of the while-loop in Line 3 attackers can employ ASICS and FPGAs to gain a computing of Algorithm 5. This is because a series of 1 branches or 2 advantage of up to 100,000x [16] over general purpose com- branches can execute in very quick succession when branch 5 puters, the cost of memory remains the same for both general is not taken inbetween them. purpose machines and ASICS/FPGAs. This makes it difficult To make sure that the trace only contains deletions, we for attackers to compute the memory-hard hash function at a calibrate our parser to be extremely conservative with adding significantly lower cost than honest users, who must compute 1s and 2s to the trace, only adding them to the trace when the the hash with general purpose computers. For this exact rea- signal is extremely clear. This ensures that only deletions, and son, scrypt is an attractive option for hashing passwords, as not additions, appear in the trace. [10] proved that scrypt is maximally memory-hard under the Attack Results. After collecting a trace with only dele- parallel random oracle model. tions in branches 1 and 2, we passed the noisy trace to the This memory-hardness property of scrypt comes at a price, branch-and-prune algorithm. Within just 34 ms, the algorithm however. Namely, [9] show that no function can be both max- found the correct key, with 18 branches inserted, and with all imally memory-hard and input oblivious; as a consequence, 703 branches correctly recovered. This key was also the first scrypt is inherently vulnerable to cache side-channel attacks, one output by the program, and after continuing to run the and in order to mitigate our attack, a compromise is required branch-and-prune program for 10 minutes, this key had the between input obliviousness and memory-hardness. largest score, making it clear that it was the correct key. Hav- As such, we recommend replacing scrypt with an alterna- ing recovered all SUBS[i] and SHIFT S[i], we then used the tive option, such as one of the side-channel resistant variants method described by [1] to trivially recover BEEA’s inputs. of Argon2 [14], the winner of the 2015 Password Hashing Implications for Chrome’s Password Leak Detection Pro- Competition. Argon2i is a variant of Argon2 that is com- tocol. A malicious server that successfully launches this pletely constant-time; however, it offers the weakest memory- attack can recover the client’s secret, a which was used in hardness of the Argon2 variants. This may be unappealing for Line 7 of Algorithm 1 to blind the hash of its credentials as Password Leak Detection, where resistance against parallel Qa . After Qa is sent to the Password Leak Detection server GPU cracking attacks is highly desirable. as part of its request, the server can easily compute a−1 and A compromising solution is Argon2id, which aims to strike −1 use it to unblind the client’s hash as (Qa )a = Q, where Q is a balance between memory-hardness and side-channel resis- the unblinded hash of the client’s credentials. tance. This is accomplished by making the first pass over This completely violates the security guarantees of the input oblivious, while the second half is input dependent, Chrome’s Password Leak Detection protocol, as it was de- thereby reducing the amount that can be learned by a side- signed to allow client’s to safely query the Password Leak channel attacker. This, however, means that Argon2id is not Detection service without having to place any trust in the completely constant-time, and still leaks some amount of server. In this case however, the client is essentially sending information to side-channel attackers. We caution against per- a hash digest of their credentials to the server, allowing the mitting any side-channel leakage at all; as we demonstrated server to run offline dictionary attacks using lists of compro- with our attack against scrypt, even an extremely limited view mised credentials aiming to breach the client’s account. of the victim’s memory accesses can potentially lead to a complete breach of security. 7 Mitigations Hash-to-Curve. In contrast, protecting the hash2curve por- The defacto standard for mitigating cache side-channel attacks tion of Password Leak Detection is comparatively simple; it in software is to make use of the constant-time programming does, however, require a slight change in protocol, due to the paradigm. In this style of programming, the control flow of current hash2curve algorithm’s usage of a rejection sampling the program must not depend in any way upon the program’s method, which is inherently non constant-time. Instead, us- input; moreover, no accessed memory address can depend ing one of the constant-time hash-to-curve implementations upon the input [42]; in other words, the execution of the described in [34] is sufficient to mitigate our attack against program must be completely oblivious to its input. the hash2curve portion of Password Leak Detection. Mitigating against the three vulnerabilities described in this Modular Inversion. Similarly, the BEEA algorithm used for paper, however, is not as easy as simply replacing vulnera- modular inversion by Chrome is inherently non constant-time; ble components with constant time implementations. 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