--- type: Whitepaper title: "Triple Handshakes and Cookie Cutters: Breaking and Fixing Authentication over TLS" description: "A man-in-the-middle can synchronise the master secret across two TLS connections by chaining RSA or Diffie-Hellman key exchange, session resumption and renegotiation, so the client's later authentication is redirected to the attacker's server. This breaks tls-unique channel binding, PEAP, SASL and channel-bound cookies, and related truncation tricks cut HTTPS headers to strip cookie flags." resource: "https://www.ieee-security.org/TC/SP2014/papers/TripleHandshakesandCookieCutters_c_BreakingandFixingAuthenticationoverTLS.pdf" tags: [whitepaper, webseclist-reference, tls, https, cookie, auth-bypass, session-fixation, formal-analysis, mitigation, owasp-a01-2021, owasp-a02-2021, owasp-a07-2021] generated: by: webseclist-refs/1 at: "2026-08-14T21:00:46+00:00" status: stable stale_after: 2027-08-14 sources: - id: original resource: "https://www.ieee-security.org/TC/SP2014/papers/TripleHandshakesandCookieCutters_c_BreakingandFixingAuthenticationoverTLS.pdf" title: "Triple Handshakes and Cookie Cutters: Breaking and Fixing Authentication over TLS" author: Karthikeyan Bhargavan, Antoine Delignat-Lavaud, Cédric Fournet, Alfredo Pironti, Pierre-Yves Strub also_at: [] authors: - Karthikeyan Bhargavan - Antoine Delignat-Lavaud - Cédric Fournet - Alfredo Pironti - Pierre-Yves Strub canonical_url: "" cited_by: - "2014.md:62" commit: "" content_sha256: 753497b94d64a956a07b14ea1db3331ba06511ef8c7c47809491dd0d4fdb82dc depth: full depth_reason: default kind: whitepaper language: "" licence: unknown original_url: "https://www.ieee-security.org/TC/SP2014/papers/TripleHandshakesandCookieCutters_c_BreakingandFixingAuthenticationoverTLS.pdf" published: "" publisher: "" publisher_english: "" raw_sha256: 1222724bd595cc533bbae6d9abdc3fe46d7311c7fd68f16a20dfd4e0f5e26f06 retrieved_from: "https://www.ieee-security.org/TC/SP2014/papers/TripleHandshakesandCookieCutters_c_BreakingandFixingAuthenticationoverTLS.pdf" retrieved_kind: stored retrieved_utc: "2026-08-14T21:00:46+00:00" slug: triple-handshakes-cookie-cutters-breaking-fixing-authentication-over-tls snapshot: "" title_english: "" translation_file: "" translation_of: "" --- # Triple Handshakes and Cookie Cutters: Breaking and Fixing Authentication over TLS **Triple Handshakes and Cookie Cutters: Breaking and Fixing Authentication over TLS** - Karthikeyan Bhargavan, Antoine Delignat-Lavaud, Cédric Fournet, Alfredo Pironti, Pierre-Yves Strub, Publisher not stated. - Published: date not stated - Original: - Preserved from: https://www.ieee-security.org/TC/SP2014/papers/TripleHandshakesandCookieCutters_c_BreakingandFixingAuthenticationoverTLS.pdf (stored) on 2026-08-14 - Licence: unknown Rights remain with the original author and publisher. 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Triple Handshakes and Cookie Cutters: Breaking and Fixing Authentication over TLS Karthikeyan Bhargavan∗ , Antoine Delignat-Lavaud∗ , Cédric Fournet† , Alfredo Pironti∗ and Pierre-Yves Strub‡ ∗ INRIA Paris-Rocquencourt † Microsoft Research ‡ IMDEA Software Institute Abstract—TLS was designed as a transparent channel abstrac- sessions, validating certificates, etc. Meanwhile, TLS appli- tion to allow developers with no cryptographic expertise to protect cations continue to rely on URLs, passwords, and cookies; their application against attackers that may control some clients, they mix secure and insecure transports; and they often ignore some servers, and may have the capability to tamper with network lower-level signals such as handshake completion, session connections. However, the security guarantees of TLS fall short resumption, and truncated connections. of those of a secure channel, leading to a variety of attacks. We show how some widespread false beliefs about these guar- Many persistent problems can be blamed on a mismatch antees can be exploited to attack popular applications and defeat between the authentication guarantees expected by the appli- several standard authentication methods that rely too naively on cation and those actually provided by TLS. To illustrate our TLS. We present new client impersonation attacks against TLS point, we list below a few myths about those guarantees, which renegotiations, wireless networks, challenge-response protocols, we debunk in this paper. Once a connection is established: and channel-bound cookies. Our attacks exploit combinations of 1) the principal at the other end cannot change; RSA and Diffie-Hellman key exchange, session resumption, and 2) the master secret is shared only between the two peers, renegotiation to bypass many recent countermeasures. We also so it can be used to derive fresh application-level keys; demonstrate new ways to exploit known weaknesses of HTTP 3) the tls-unique channel binding [6] uniquely identi- over TLS. We investigate the root causes for these attacks and propose new countermeasures. At the protocol level, we design fies the connection; and implement two new TLS extensions that strengthen the 4) the connection authenticates the whole data stream, so it authentication guarantees of the handshake. At the application is safe to start processing application data as it arrives. level, we develop an exemplary HTTPS client library that The first is widely believed to be ensured by the TLS renego- implements several mitigations, on top of a previously verified tiation extension [49]. The second and third are used for man- TLS implementation, and verify that their composition provides in-the-middle protections in tunneled protocols like PEAP and strong, simple application security. some authentication modes in SASL and GSS-API. The fourth forms the basis of HTTPS sessions on the web. I. T RANSPARENT T RANSPORT L AYER S ECURITY ? These assumptions are false, and this enables various at- TLS is the main Internet Standard for secure communica- tacks, even against applications using the latest, fully-patched tions and still, after 20 years of practice, the security it provides TLS 1.2 implementations. Whether these attacks should be to applications remains problematic. blamed on the protocol or its usage, we argue that the transport and application protocols must be analyzed together to achieve I-A A PPLICATIONS VS P ROTOCOLS . By design, TLS intends reliable, meaningful, application-level security. to provide a drop-in replacement of the basic networking On the other hand, our paper does not challenge the cryp- functions, such as connect, accept, read and write, tographic security of the core constructions of TLS—most that can effortlessly protect any application against a net- of our attacks apply even under the (theoretical) assumption work attacker without the need to understand the protocol or that clients and servers only use cryptographically strong its underlying cryptography. Pragmatically, TLS offers much ciphersuites, as formalized, for example, in [15, 35, 29, 16]. flexibility, so the security properties provided by the protocol [43, 35, 32, 29] and its implementations [20, 14, 15] depend I-B N EW ATTACKS OVER TLS. We report new practical at- on how TLS is used. For instance, if the application enables an tacks against applications that rely on TLS for their security. unsuitable ciphersuite [4], uses compression [25], or ignores The first family of attacks uses a combination of successive state changes [45], it opens itself to attacks. Furthermore, TLS handshakes over multiple connections to disrupt client applications-level security mechanisms increasingly seek to authentication. The second family of attacks targets HTTPS benefit from the underlying TLS connection by reusing its message integrity but may apply to other application protocols. authenticated peer identities, key materials [48], and unique Triple Handshakes Considered Harmful (§V, §VI) We identifiers [6]. first point out unknown key-share [17] vulnerabilities in RSA, As a consequence, TLS libraries provide low-level APIs DHE, and abbreviated handshakes, and we compose them that expose many details of the cryptographic mechanisms to implement a malicious TLS proxy that can synchronize and certificates negotiated during successive handshakes. Some the keys on separate connections with honest peers. These application-level libraries, such as CURL, seek to recover the vulnerabilities do not in themselves constitute attacks on the simplicity of a secure channel by implementing an abstraction integrity and confidentiality guarantees of TLS. However, we layer that smooths over the details of TLS by managing show that they enable new man-in-the-middle attacks that break a variety of authentication mechanisms built over TLS, required by modern browsers and web services. However, it including (a) client-authenticated TLS renegotiation—for ex- automatically handles all the details of the underlying TLS ample if a client presents her certificate to two TLS servers, connections, including multiple handshakes, resumption and one can impersonate the client at the other; (b) compound negotiation, and truncations. authentication in tunneled protocols; (c) channel bindings for I-D M AIN C ONTRIBUTIONS . We describe a new class of application-level challenge-response protocols; and (d) channel man-in-the-middle attacks against authentication over TLS, bindings for bearer tokens. We report concrete attacks against targeting the resumption and renegotiation features of the published specifications and popular applications in all these handshake. We also present new exploits on HTTPS sessions categories, including mainstream browsers and HTTP client based on cookie-forcing and truncation. We apply these attacks libraries, VPN applications, wireless applications, and mail and to break the expected authentication guarantees of several chat servers. state-of-the-art protocols, libraries, applications, and web ser- Truncating Headers & Forcing Cookies (§III) Indepen- vices. We have contacted many vendors, proposing immediate dently, we show that web browsers and servers often ignore mitigations and countermeasures, as well as more long-term TLS disconnections and tolerate ill-formed messages, thereby fixes to the corresponding protocol specifications. Our TLS- enabling message truncations. Although this vulnerability is level proposals are consolidated in patches for OpenSSL and generally known [13, 52], we show how to apply truncation miTLS. We have also built and verified a basic high-level to HTTP headers and HTML forms, opening new exploits. In HTTPS API on top of miTLS, to validate our main application- particular, our attacks completely defeat cookie-based authen- level recommendations in a simplified setting. tication. We also show new exploits based on known attack vectors like cookie-forcing and its use for login CSRF [12, 18]. Contents §II reviews the dangers of application security over In particular, we show that building new application-level TLS. §III illustrates these dangers by presenting new attacks protocols such as single sign-on and synchronization protocols caused by truncating HTTPS traffic and forcing cookies. §IV using cookies is foolhardy; they amplify login CSRF attacks recalls the relevant protocol aspects of TLS. §V describes a ma- and enable network attackers to steal users’ private files. licious TLS proxy that synchronizes connections between TLS clients and servers. §VI presents new proxy-based attacks on I-C T OWARDS V ERIFIED A PPLICATION S ECURITY. In light applications that use client authentication. §VII discusses TLS of the two families of attacks outlined above, how to ensure countermeasures, implemented in OpenSSL and miTLS. §VIII that a TLS application properly handles its interactions with illustrates application-level countermeasures, demonstrating a the TLS API? How to reliably lift TLS security to application simple, provably secure HTTPS API on top of miTLS. §IX security? Broadly, we can either build countermeasures into discusses impact, limitations and responsible disclosure of the TLS; or carefully implement and verify simpler security APIs attacks presented in this paper. over TLS; or, less robustly, promote prudent practices for writing secure applications over TLS. Online Materials An extended version of this paper, the two patches implementing our proposed countermeasures for Proposed TLS Extensions (§VII) One approach is to OpenSSL and for miTLS, our verified implementation of strengthen the protocol to provide more robust security. To miHTTPS and further experimental data are available online this end, we propose two new TLS extensions that prevent the at https://secure-resumption.com. attacks of §VI without the need to change applications. These extensions have a negligible impact on performance and code complexity, and can be deployed while preserving backward II. TLS I NTERFACES AND THEIR S AFE U SAGE compatibility. They apply to all protocol versions from SSL3 to TLS 1.2, as well as DTLS. To validate them experimentally, Modern clients and servers interact with TLS in ways far we implemented and tested patches for two existing TLS beyond the original intended interface. We discuss typical implementations: OpenSSL and miTLS. As future work, we usages of the protocol, relevant to the attacks of §III and §VI. plan to formally model their security benefits by extending II-A S ESSION AND C ERTIFICATE M ANAGEMENT. HTTP is the verified cryptographic model of miTLS [15, 16]. by far the most widely used application protocol over TLS. Simple Verified HTTPS over TLS (§VIII) In principle, Even the most basic HTTP operation, getting a file from carefully-written applications can defend against these attacks, a given URL, may require multiple connections to different without the need to change TLS. To validate our main recom- servers due to redirections, authentication requests, temporary mendations, and show that “transparent” application security errors, and many other factors. Thus, any HTTPS client can indeed be achieved over TLS, we program miHTTPS: must manage and isolate multiple TLS sessions with different a simple HTTPS library in F#, similar to CURL, on top principals: if a client ever uses a cached session with the wrong of miTLS. We specify its intended security properties and server, the security guarantees of TLS collapse. we verify them using F7, a type-based verification tool. Similarly, any TLS application must implement a server Thus, we formally relate the precise, low-level TLS API certificate validation procedure, which can combine subject provided by miTLS to a simpler, more abstract HTTPS API. name and certificate purpose validation, pinning of certification In combination, we obtain the first cryptographically-verified authorities, trust on first use (TOFU), among others [22, 28]. implementation for HTTPS. In its current state, miHTTPS is Once again, any error in this process may completely void the a proof-of-concept: it does not nearly provide the flexibility security guarantees of TLS. While session and certificate management are critical to To address the shortcomings of authentication at the applica- the security of the protocol, they are implemented at the tion level, new solutions have been recently proposed to expose application level in contradiction to the network abstraction of values taken from the TLS handshake to applications in order TLS. Even when TLS libraries provide default functionality for to bind their bearer tokens and challenge-response protocols these operations, they are not necessarily secure; for instance, to the underlying TLS channel. Hence, tunneled wireless pro- OpenSSL shares the client-side session cache between all tocols like PEAP [42] use compound authentication schemes connections, even if they are to different hosts, unless it is [44] to protect against rogue access points. SASL mechanisms explicitly partitioned by the application. like SCRAM [39] use TLS channel bindings [6], in partic- II-B E XPOSURE TO TLS E VENTS . Another recurrent prob- ular the tls-unique binding, to prevent man-in-the-middle lem with TLS APIs is the way they should expose transport- attacks even on anonymous TLS connections. Channel ID [10], level events to the application. In this paper, we focus on two a follow up to Origin-Bound Certificates [24], proposes that the events that can lead to attacks if ignored by the application: client generate a long-lived pair of keys associated with each renegotiation, and TCP connection closure. top-level domain it connects to. The public key is treated as a client identifier and, by binding bearer tokens such as cookies Once a TLS connection is established, most applications to this public key, the server can ensure they can only be used typically only use read, write and close. How can a by the client they have been issued for, thus mitigating token TLS library notify the application when renegotiation occurs? compromise. §VI studies the assumptions such mechanisms What if the cipher or the peer certificate changes? At best, make about TLS and presents attacks on a number of them. the read primitive can return a non-fatal error code (like in GnuTLS) which the application can either ignore or use to enforce further checks on the new parameters. At worst, the III. T RANSPORT- LAYER ATTACKS ON HTTPS change is only visible if the application keeps polling specific As a case study of the API problems of §II, we consider session parameters. To protect applications that ignore such the use of HTTP over TLS [47]. In HTTP, messages consist events from man-in-the-middle attacks [45], most TLS libraries of two parts: the headers and an optional body, separated by implement a protocol extension [49]. §VI-A shows how these an empty line. Headers consist of colon-separated name-value applications can still be attacked despite this countermeasure. pairs, each terminated by a line break. The first header line Since SSL3, the closure of a connection must be notified is special: in requests, it contains the method (either GET or to the other party with an authenticated protocol alert called POST), path, and protocol version; in responses, it contains close_notify. Without this graceful closure, a man-in-the- the protocol version, status code, and status message. The middle may have closed the TCP connection in the middle HTTP body is formatted according to the headers: by default, of a TLS connection. To make this distinction, TLS libraries its length is specified in the Content-Length header; should return a special error code when truncation is detected, if the Content-Transfer-Encoding header is set to signaling to the application not to process any partial data chunked, the body is a sequence of fragments, each prefixed that may be buffered. However, in several implementations, the by the fragment length, terminated by an empty fragment. read primitive returns the number of bytes read, while error Due to the variety of (not necessarily correct) HTTP im- checking requires manual verification of a different parameter. plementations, most clients are very permissive when parsing Many applications do not distinguish between normal and HTTP. For instance, they often accept message bodies whose unexpected closure, sometimes deliberately for compatibility. length does not match the one indicated in the headers, or Another class of problems appears when TLS is an optional missing the last empty fragment in the chunked encoding. feature of the application protocol, or if state is shared between For authentication, almost all websites rely on cookies, encrypted and plaintext connections. §III illustrates how to which are name-value pairs set by servers in the Set-Cookie exploit these issues against HTTP. header and sent back by clients in the Cookie header of subsequent requests. The cookie store is shared between HTTP II-C C LIENT AUTHENTICATION . Applications can use vari- and HTTPS connections, opening up a variety of attacks. ous mechanisms for client authentication: client certificates (e.g. in browsers, for virtual private networks, and for wireless III-A C OOKIE I NTEGRITY. Modern web security policies are access points), bearer tokens (e.g. HTTP sessions cookies and expressed in terms of origin, i.e., the combination of protocol, OAuth access tokens), or challenge-responses protocols (e.g. domain and port. Hence, HTTP requests and JavaScript inter- HTTP digest authentication, and several SASL mechanisms actions are unrestricted within the same origin, and strictly used by mail and chat servers). regulated across different origins [57]. In contrast, cookie TLS client authentication is generally considered the safest, policies rely on domain and path; furthermore, cookies may but is seldom used. Weaker mechanisms that rely on bearer be set for any domain suffix and path prefix of the current tokens are more common, but they allow complete long-term page, e.g. http://y.x.com/a can set cookies with domain impersonation of the user when a token is compromised. x.com and path /. This discrepancy causes major problems: Challenge-response authentication within TLS tunnels offers • Protocol: since there is no separation between HTTP and better protection, but is still vulnerable to man-in-the-middle HTTPS, by default, cookies set on encrypted connec- attacks [8, 41]: if the user is willing to authenticate on a server tions are also attached to plaintext requests, in plain sight controlled by the attacker, the attacker can forward a challenge of the attacker. To prevent this, the secure flag can be from a different server to impersonate the user at that server. sent when setting the cookie to indicate to the browser never to send this cookie unencrypted. This protects the TABLE I. TLS T RUNCATION IN B ROWSERS confidentiality of cookies, but not their integrity, as it In-Header Content-Length Missing last chunked still possible to overwrite secure cookies over HTTP. truncation ignored fragment ignored • Domain: domains prefixed with a dot will match any Android 4.2.2 Browser Android Chrome 27 subdomain. Thus, a request to a.x.com attaches cook- Android Chrome 28 ies set for .x.com, but not those set for b.x.com. A Android Firefox 24 page may set cookies on any of its own domain suffix Safari Mobile 7.0.2 Opera Classic 12.1 that is not a public (such as “com” or “co.uk”), leading Internet Explorer 10 to related-domain attacks. • Port: since the port number is ignored, and even if a website is only served over TLS, an attacker can still for instance after a successful login, the new URL given in the use some unencrypted port to tamper with its cookies. Location header typically includes parameters taken from the request (e.g., the page the user was trying to access before Cookies with the same name but different domain or path are logging in). Such parameters are often under attacker control, stored separately; all matching cookies are sent back in the and allow targeted truncation in response headers as well. Cookie header in an unspecified order. Finally, there is a limit on the number of cookies that can be stored for each top-level Truncating Responses Recall that browsers do not at- domain name (e.g. x.co.uk). Beyond this limit, typically tach cookies set with the secure flag to HTTP requests. around 1000, older cookies are automatically deleted. Thus, In the Set-Cookie header, however, the flag occurs af- an attacker can reliably delete legitimately set cookies. ter the cookie, so the attacker can selectively truncate it Cookie forcing, cookie fixation, and cookie tossing all refer and redirect the user to an unencrypted URL to recover to tampering with cookies, either from the network or from the cookie value. Concretely, consider a login form at a related subdomain. These issues have been well known for https://x.com/login?go=P that sets a session cookie years, and many proposals address them [12, 18, 24], but there and redirects the user to https://x.com/P. The headers is still no way to defend against cookie forcing by a network of the response are as follows: attacker that works on all current browsers. Experimentally, we were able to force sessions on the top 10 Alexa websites HTTP/1.1 302 Redirect in the US, despite the mitigations deployed on some of them. Location: https://x.com/P Set-Cookie: SID=[AuthenticationToken]; secure Worse, the impact of such forcing attacks has increased Content-Length: 0 considerably recently. For instance, many websites rely on single sign-on services for authentication. If the session on The attacker can chose P such that the first TLS fragment ends the identity provider (such as Facebook, Twitter or Google) just before ‘;’ and close the connection before the second is replaced with the attacker’s, the victim may unwittingly fragment is sent, allowing the cookie to be stored without the associate his accounts on many websites with the attacker’s secure flag (and thus, visible to the attacker over HTTP). We identity, even after leaving the attacker’s network. Furthermore, successfully mounted this attack against Google Accounts. in modern websites, many operations are performed asyn- The attack is possible because some browsers, including chronously. Thus, if a session is forced onto the browser before Chrome, Opera, and Safari, accepted incomplete HTTP re- such an action, it may be associated with the attacker’s account sponses (missing an empty line at the end of headers). We without any feedback to the user. Finally, some browsers rely reported the vulnerability to each vendor; their responses are on web login forms to provide features such as synchronization given in §IX-A. Table I summarizes the possible truncations in of tabs, bookmarks and stored passwords. We found that login current browsers; we focus on mobile versions because they are CSRF attacks could trigger such features; even though a user more likely to connect to untrusted networks. While header- confirmation dialog is shown with the account name of the truncation attacks have mostly been fixed, chunked-body- forced session, it provides a tempting phishing target. truncation attacks remain possible on HTML and JavaScript. III-B T HE C OOKIE C UTTER ATTACK . As discussed in §II, Truncating Requests While most servers do not accept most HTTP software does not enforce proper TLS termination, truncated headers, some do accept a truncated body. In the letting the attacker truncate a message at any TLS-fragment case of POST requests, typically used when submitting a form, boundary by closing the underlying TCP connection. If the the parameters are sent in the body of the request. This is attacker controls the length of some of the contents of the most notably the case of requests sent through Apache SAPI message, he may chose a specific truncation point. Although modules, such as PHP. The main difficulty when truncating a this pattern has been exploited before to delete entire HTTP re- POST request is to guess the length of the body parameters, quests or to truncate message bodies [13, 52], we demonstrate which may be difficult since they often contain user input. new truncation attacks within headers of HTTP messages. Consider a scenario where the victim invites one of her A network attacker can trigger a request with any path and friend bob@domain.com on a social network where the parameters (in fact, any website can trigger such requests to attacker wants to access her profile. The attacker registers the any other website) and inject data into its Cookie header us- domain domain.co and monitors the victim as she accesses ing forcing techniques, thus controlling the TLS fragmentation the invitation page (for instance, by inspecting the length of of the request. In response headers, when a redirection occurs, the returned page). The query to truncate is of the form: POST /invite.php HTTP/1.1 Host: socialnetwork.com Content-Type: application/x-www-form-urlencoded Cookie: SID=X; ForcedByAttacker=Z Content-Length: 64 csrf_token=Y&invite=bob@domain.com When the query is sent, the attacker truncates it such that the invitation will be sent to bob@domain.co. The victim gets a blank page due to the truncation, and may try the request again. Meanwhile, the attacker receives credentials to access the victim’s profile. We were able to mount this attack on a popular social network that uses Apache and PHP. III-C TLS C ONNECTION I NTEGRITY. Because most users connect to websites using plain HTTP, even if a website redirects all unencrypted connections to HTTPS, it is easy for a man in the middle to forward HTTPS contents over HTTP to the user, rewriting all links and pointers to encrypted pages. Fig. 1. The TLS Handshake This attack, called SSL stripping [37], is very popular thanks to simple tools to mount it on public wireless networks. attacker can re-enable SSL stripping, cookie forcing, and the To protect against SSL stripping, several browsers support cookie secure flag truncation attack above even on websites HTTP Strict Transport Security [30] (HSTS), which introduces that enable HSTS, defeating the purpose of this standard. a Strict-Transport-Security header for websites to For websites that do not deploy HSTS, browser extensions indicate that the browser should always connect to its domain have been developed to force the use of HTTPS on a given over TLS, regardless of the port. The header includes a list of websites. However, it is worth noting that such ad max-age value, specifying how long this indication should hoc mechanisms have their own flaws. For example, HTTPS be enforced, and an optional includeSubDomains flag, Everywhere [2] allows HTTP connections when the server port indicating that the policy also applies to all subdomains. is non-standard. Cookie policies ignore the port number, so HSTS has several known weaknesses. The first problem various attacks like cookie forcing remain possible. is bootstrapping: the user may use HTTP the first time it connects to the website, before receiving the HSTS header in IV. TLS P ROTOCOL : C ONNECTIONS , S ESSIONS , E POCHS the response. This bootstrapping problem is typically mitigated by browsers that use a pre-registered HSTS domain list for The TLS protocol is commonly used over TCP connections sensitive websites that wish to opt-in to this feature. to provide confidentiality and integrity for the bytestreams Second, HSTS preserves cookie integrity only when enabled exchanged between a client (C) and a server (S). Next, we on the top level domain with the includeSubDomains flag, recall the main subprotocols of TLS and the attacks directly and if the user visits this domain first [18]. This is an expensive relevant to this paper. (The online version discusses other prior requirement for large websites, as it forces all contents for attacks on handshake integrity.) We assume some familiarity the entire domain to be served over HTTPS. We found that with TLS; we refer to the standard [23] for the details and to not a single website from the top 10,000 Alexa list is using other papers for a discussion of previous proofs [35, 43] and the includeSubDomains option on their top-level domain, attacks [40, 22]. even though some are indeed using HSTS. Thus, in practice, IV-A F ULL H ANDSHAKE . Once a TCP connection has been HSTS is not used to prevent cookie forcing attacks. established between a client and a server, the TLS handshake We found a new attack to bypass HSTS on some clients. protocol begins. The goals of the handshake are to authenticate A network attacker can truncate the Strict-Transport- the server and (optionally) the client; to negotiate protocol Security header after the first digit of the max-age pa- versions, ciphersuites, and extensions; to derive authenticated rameter. If the client accepts and processes this header, the encryption keys for the connection; and to ensure agreement HSTS entry for that website will expire after at most ten on all negotiated parameters. (A ciphersuite selects a key seconds, after which HTTP connections to the domain will exchange mechanism KEX ALG for the handshake and an be allowed again, even if the domain has pre-registered to the authenticated encryption mechanism ENC ALG for the record HSTS domain list on the browser. protocol.) Concretely, to attack x.com, the man-in-the-middle takes Figure 1 shows the full handshake with mutual authenti- any HTTP request for any server and redirects it to a page on cation. First, the client sends a client hello message with a x.com that returns a parameter-dependent Location header maximum protocol version pvmax , a random nonce cr, and followed by the Strict-Transport-Security header. a set of proposed ciphersuites and extensions. The server We successfully tested the attack on Chrome, Opera, and chooses a version pv, a ciphersuite, and a subset of these Safari. We further note that by using this attack first, a network extensions, and responds with its own nonce sr and a session identifier sid . The server then sends its X.509 certificate chain cert S and public key pk S . Depending on KEX ALG, it may send additional key materials in a key exchange message kex S . It may also send a certificate request message if it requires client authentication. The client responds with its own certificate chain cert C and public key pk C (if required), followed by its own key exchange message kex C . If the client sends its certificate, it also sends a signed hash sig C of the current log (log 1−8 , obtained by concatenating messages 1–8) in a certificate verify message. At this point in the protocol, both the client and the server Fig. 2. Abbreviated TLS Handshake can compute a shared pre-master secret pms from kex C and kex S , then use pms along with the nonces to derive a master accept named curves within a fixed set, whereas DHE allows secret ms, and use ms to derive keys for the connection and to the server to choose arbitrary DH group parameters. verify the handshake integrity; these computations are detailed Other key exchanges are less common on the web but below. To complete the handshake, the client signals a change useful in other applications. In TLS-PSK, the client and server of keys with a change cipher spec (CCS) message followed authenticate one another using a pre-shared key instead of cer- by a finished message that contains the client verify data cvd tificates. In TLS-SRP, the client uses a low-entropy password obtained by MACing the current handshake log (log 1−9 ) with instead of a certificate. In DH anon, both client and server key ms. Similarly, the server sends its own CCS and a finished remain anonymous, so the connection is protected from passive message that contains the server verify data svd , obtained by eavesdroppers but not from man-in-the-middle attackers. MACing the whole handshake log 1−9,11 . (The CCS messages are not included in the logs.) IV-B T HE R ECORD P ROTOCOL . Once established, a TLS When the client is not authenticated, messages 5, 7, 9 connection provides two independent channels, one in each are omitted. When the server does not contribute to the key direction; the record protocol protects data on these two exchange, e.g. with RSA, message 4 is omitted. channels, using the authenticated-encryption scheme and keys RSA Handshake If the key exchange in the negotiated provided by the handshake. Application data is split into a ciphersuite is RSA, the calculations go as follows, where stream of fragments that are delivered in-order. There is no log 1−8 is the log before message 9, log 1−9 is the log before correlation (at the TLS level) between the two directions. message 11, and log 1−9,11 is the log before message 13. (The When the client or server wishes to terminate the connection, it server key exchange value kex S is not used.) sends a close_notify alert to signal the end of its writing stream, and it may wait for the peer’s close_notify before pms = [pvmax ]|[46 bytes randomly generated by C] closing the connection. If both peers perform this graceful sig C = signed(sk C , log 1−8 ) closure, they can both be sure that they received all data. kex C = rsa(pk S , pms) However, this is seldom the case in practice. There are several attacks on the confidentiality of the record ms = prf(pms, “master secret”, cr|sr) protocol [e.g. 5]; attacks on integrity are less common [e.g. 15]. keys = prf(ms, “key expansion”, sr|cr) IV-C S ESSION R ESUMPTION . Full handshakes involve mul- cvd = prf(ms, “client finished”, hash(log1−9 )) tiple round-trips, public key operations, and (possibly) svd = prf(ms, “server finished”, hash(log1−9,11 )) certificate-revocation checks, increasing latency and server load [53]. In addition, abbreviated handshakes enable clients DHE Handshake If the negotiated key exchange is ephemeral and servers that have already established a session to quickly Diffie-Hellman (DHE), then S chooses group parameters (p, g) set up new connections. Instead of establishing a new master and a fresh key pair (KS , g KS ); it sends (p, g, g KS ) in kex S , secret, both parties reuse the master secret from that recent signed along with cr and sr with its private key sk S . The session (cached on both ends), as shown in Figure 2. client generates its own key pair (KC , g KC ) and responds with The format of the cached session data depends on the kex C = g KC . Both parties compute pms = g KC ∗KS . The rest TLS implementation, but [50] recommends that it contains of the computations are the same. at least the master secret, protocol version, ciphersuite, and kex S = signed(sk S , cr|sr|p|g|g KS mod p) compression method, along with any certificate used. The client sends a client hello, requesting the server to kex C = g KC mod p resume the session sid , with a new client nonce cr . If the pms = g KC ∗KS mod p (with leading 0s stripped) server has cached this session, it may then respond with a server hello with a new server nonce sr and the same sid Other variations Besides RSA and DHE, mainstream TLS and algorithms as in the initial handshake. The server then implementations support variations of the Diffie-Hellman key immediately sends its CCS and finished message, computed exchange implemented using elliptic curves. The handshake as a MAC for the abbreviated handshake log. The client for these is similar to DHE, but with some notable dif- responds with its own CCS and finished message, computed ferences. For example, most ECDHE implementations only as a MAC of the whole resumption log. The computation of keys and verify data are as follows, where log 1−2 consists of In response to this attack, a new ‘mandatory’ TLS extension the messages 1’ and 2’, while log 1−2,4 includes 1’, 2’ and 4’: has been proposed and deployed for all versions of TLS [49]. This extension includes the verify data of the previous hand- ms = [cached for (S, sid )] shake within the client and server hello messages of the keys = prf(ms, “key expansion”, sr |cr ) renegotiation handshake, thereby cryptographically binding the two handshakes (and, recursively, any preceding handshake svd = prf(ms, “server finished , hash(log1−2 )) on the same connection). As a result, as each handshake cvd = prf(ms, “client finished , hash(log1−2,4 )) completes, both peers can be confident that they agree on all The completion of an abbreviated handshake implicitly epochs on their connection. Informally, the principals at each confirms to each participant that they share the same session endpoint must remain the same, even if the certificates change. master secret. Hence, if both peers are honest, they must As shown in §V, this countermeasure still does not suffice have matching session parameters—those negotiated in the to eliminate renegotiation attacks across several connections. initial handshake. Because of its efficiency, resumption is IV-E I MPLEMENTATIONS AND API S . There are several pop- aggressively used on TLS connections. It is supported by ular implementations of TLS, including OpenSSL, GnuTLS, default in all major web browsers and web servers. A recent NSS, JSSE, and SChannel. Here, we briefly discuss the miTLS TLS extension enables servers to store their cached sessions verified reference implementation [15], whose API is distinc- at the client within encrypted tickets [50]; this mechanism tive in the detailed connection information that it offers to its makes it possible for clients to maintain long-lived sessions applications. As such, miTLS is an ideal experimental tool on with stateless server farms, at little cost to the servers. which to evaluate attacks and implement countermeasures. We use the term session resumption when the same TLS The miTLS API consists of functions to initiate and accept session is used on multiple connections, but the abbreviated connections, send and receive data, and instigate session re- handshake may also be used on an existing TLS connection to sumption, re-keying, and renegotiation. Each of these functions refresh keys and reset sequence numbers. At the end of each returns a connection handle and a ConnectionInfo structure, handshake, we say that the connection enters a new epoch. which details the current epoch in each direction (they can IV-D R ENEGOTIATION : C HANGING E POCHS . A client or a differ). For each epoch, it includes the nonces and verify data server may request a new handshake on an established TLS and points to a SessionInfo structure with the epoch’s session connection, e.g. to renegotiate the session parameters. The parameters (including ciphersuites and peer identities). It also handshake proceeds as described above, except that its mes- points to the previous epochs on the connection (if any). sages are exchanged on the encrypted TLS connection. When The API encodes the security assumptions and guarantees the handshake completes, both parties share a new session, of TLS as pre- and post-conditions on the connection state. and their connection enters a new epoch, switching to the keys The application cannot send or receive data unless the con- derived from the new session. nection is in the Open state, which means that a handshake There are many reasons why an application may want to has successfully completed with an authorized peer. When renegotiate a TLS session when it already has a working TLS a handshake completes at an endpoint, the API guarantees connection. The first is client authentication. On some servers, that, if all the principals mentioned in the ConnectionInfo client authentication is required only when accessing protected are honest, then there is exactly one other endpoint that has resources. For instance, Apache triggers renegotiation and a matching ConnectionInfo and keys. Every application data requires a client certificate on first access to a protected fragment sent or received is indexed by the epoch it was sent directory. This design improves user experience and helps on, which means that miTLS will never confuse or concatenate protect privacy by requesting authentication only when needed, two data fragments that were received on different epochs; it and prevents the client certificate being sent in the clear during is left to the application to decide whether to combine them. If the initial handshake. Other reasons may be to upgrade the the connection uses the renegotiation indication extension, the ciphersuite or replace an expiring certificate [49, §5]. Even in application gets an additional guarantee that the new epoch is this case, the server may need to provide a new certificate that linked to the old epoch. If at any point in a connection, miTLS supports, say, ECDSA signing instead of RSA. Consequently, receives a fatal alert or raises an error, the connection is no in many renegotiations, the client and server certificates and longer usable for reading or writing data. If the connection is identities after renegotiation may differ from those of the pre- gracefully closed, miTLS guarantees that each endpoint has vious handshake. Without additional protections, such identity received the entire data stream sent by its peer. Otherwise, it changes can lead to impersonation attacks. only guarantees that a prefix of the stream has been received. Renegotiation Attack Protecting the renegotiation under the V. A M AN -I N -T HE -M IDDLE TLS P ROXY S ERVER keys of the previous handshake is not enough to prevent man- in-the-middle attacks. An active network attacker can intercept We consider the following scenario. Suppose an honest TLS an initial handshake from a client to a server and forward it as client C connects to a TLS server A that is controlled by the a renegotiation within an existing TLS connection between attacker. A then connects to an honest TLS server S, and acts the attacker and the server. As a result, any data that the as a man-in-the-middle proxy between C and S, ferrying data attacker sent before the renegotiation gets attributed to the between C and S across the two independent connections. Of client, leading to a powerful impersonation attack [45]. course, A can still read and tamper with selected fragments. Now, suppose that A establishes the same keys on both TLS connections. We will show in this section how A can achieve this. Then A does not have to decrypt and reencrypt traffic between the two connections and may instead step out of the way, allowing C and S to talk directly to one another, making A’s intervention difficult to detect even with sophisticated timing measurements [9]. On its own, the scenario above does not constitute a serious attack on either connection, since both C and S are aware that they are connected to A. However, the ability of A to synchronize keys across two connections can be a stepping stone towards more dangerous attacks, as we will show in §VI. In the cryptographic key-exchange literature, this kind of key synchronization is called an unknown key-share attack [17, 34], whereby two honest parties share a key but one of them does not realize with whom it shares its key; their mutual belief in the shared secret is violated [54]. In Abadi’s terminology [3], these attacks do not disrupt any access control goals based on responsibility, but they enable an attacker to take credit for an honest principal’s message. So, if the application that uses the protocol does not reliably confirm both peers’ identities, impersonation attacks may appear [36]. In the rest of this section, we show how a malicious server A can synchronize TLS keys with C and S. To build this malicious server, we exploit three independent weaknesses in the RSA handshake, the DHE handshake, and the abbreviated handshake. We do not make any assumption about application behavior, and use only standard mechanisms implemented by mainstream TLS libraries. V-A S YNCHRONIZING RSA. Suppose C sends a client hello to A offering an RSA ciphersuite. A then forwards the client hello to S. When S responds with the server hello, A forwards it to C. Hence, the client and server nonces cr , sr and the session identifier sid are the same for both connections. Next, when S sends its certificate cert S to A, A instead sends its own certificate cert A to C. Now, C generates a pre- master secret pms, encrypts it under pk A , and sends it to A. Fig. 3. Triple handshake attack by a malicious server on client-authenticated A decrypts pms, re-encrypts it under pk S , and sends it to S. TLS renegotiation: (1) RSA/DHE full handshake, (2) abbreviated handshake Hence, both connections have the same pms and (since the for session resumption, (3) secure (RFC 5746 [49]) renegotiation handshake nonces are equal) the same master secret and connection keys, all of which are now shared between C, S, and A. Finally, C and S support RSA but prefer a different key exchange, say A completes the handshake on both connections, using ms to ECDHE, A can still force them both to use RSA by offering compute correct verify data. The messages tampered by A are only RSA in its client and server hellos. illustrated in Figure 3 (Connection 1). The RSA key exchange does not ensure different keys on At this point, C and S cache the same session that they both different connections, and there is no standard mitigations associate with A (as represented by cert A on C, and optionally, that implementations can employ to prevent it. This behavior A’s client certificate on S). The new epochs on the two would not surprise a cryptographer or protocol expert, since connections are distinguishable only by the client and server only C contributes to the key exchange. However, it is only verify data, which differ on the two connections. However, occasionally mentioned in protocol specifications [48, §5] and messages from one connection can be freely forwarded to the continues to surprise protocol designers. As shown in §VI, such other, since the keys match. Consequently, if A stepped out of connection synchronizations can defeat the man-in-the-middle the way, C and S can continue exchanging messages without protection used in tunneled protocols like PEAP. realizing that the principal on the other end has changed. V-B S YNCHRONIZING DHE. Suppose that C (or S) refuses Variants and Mitigations The above trace is robust to vari- RSA ciphersuites, but accepts some DHE ciphersuite. We show ations in the key exchange. If S demands a client certificate, that A can still synchronize the two connections, because A can provide its own certificate, and this does not affect the the DHE key exchange allows the server to pick and sign synchronization of the master secret or connection keys. If both arbitrary Diffie-Hellman group parameters, and any client that accepts the server certificate and signature implicitly trusts same attack, but all the TLS implementations we tested only those parameters. support well-known named curves standardized by NIST. In this scenario, A substitutes its own certificate for S’s (as V-C S YNCHRONIZING A BBREVIATED H ANDSHAKES . Sup- with RSA), then changes the Diffie-Hellman group parameters pose C, A, and S have synchronized sessions and connections, in the server key exchange message, and finally changes the as described above. If C attempts to resume the session with A client’s public key in the client key exchange message. over a new connection, A can then synchronize this new Suppose S offers a prime p, generator g, and public key connection with a new connection to S. In fact, abbreviated PS = g KS mod p. A replaces p with the non-prime value p = handshakes are easier to synchronize than full handshakes. PS (PS − 1) and signs the parameters with its own private key. When C sends its client hello requesting session resumption When C sends its own key exchange message with public key on a new connection, A simply forwards the request to S, and PC = g KC mod p , the attacker replaces it with the public key forwards S’s response to C unchanged. C and S complete g and sends it to S. Our choice of p ensures that PS has order the handshake through A, re-using the master secret known 1 in the group Zp∗ , or equivalently ∀x > 0, PSx = PS mod p . to C, S, and A, as shown in the top half of Connection 2 in Other values of the form p = q(PS −1) also lead to PS having Figure 3. The resulting epochs on the two connections have the a low order in Zp∗ . Upon receiving this message, C computes same keys, also shared with A. The new epochs are, in fact, pms = PSKC mod PS (PS − 1) more synchronized than the epochs on the original connection: the client and server verify data on these epochs are also the = PS mod PS (PS − 1) same. Hence, after resumption, the only noticeable difference = PS (with leading 0s stripped) between the two connections is that the C-A connection has a session with server identity cert A while the A-S connection while S computes pms = g KS mod p = PS . Finally, both has a session with server identity cert S . All other differences connections share the same pms, ms, and derived keys. have been erased. This is important for the attacks in §VI. Variants and Mitigations The authenticated Diffie-Hellman The ease with which resumed sessions can be synchronized key exchange is not intrinsically vulnerable to a man-in-the- exposes the weak authentication guarantees of the abbreviated middle, as long as both parties use the same, well chosen handshake. It only ensures that the client and server share group. The key to this attack is that the attacker is able to make the same master secret, whereas applications may (and do) C accept a group with a non-prime order. In fact, p above is assume that they share the same session, which we show is not always even (and may cause errors with implementations that the case. To obtain stronger guarantees from this handshake, rely on Montgomery reduction for modular exponentiation) but in §VII we propose a TLS extension, similar to [49], that links it is easy to find odd non-primes that work just as well. the resumption handshake to the original session. The attack fails if C checks that p is prime. Yet, none of the mainstream TLS implementations perform a full primality check because it is deemed too expensive. A probabilistic VI. ATTACKS ON C LIENT AUTHENTICATION OVER TLS primality check could help, but may not guarantee that the TLS is most commonly used in the anonymous-client mode, attacker cannot find a p that defeats it. An alternative mitiga- where only the server is authenticated. Consequently, appli- tion would be to standardize a few known good Diffie-Hellman cations often deploy their own mechanisms and protocols to groups for use in TLS. Indeed, this is the approach taken in authenticate users after the TLS handshake has finished. protocols like IKEv2 and in TLS variants like SRP. Previous work shows that layering a client authentication Even when clients and servers use known groups, care must protocol within a server-authenticated secure channel is vul- be taken to validate the public key received from the peer. nerable to generic man-in-the-middle attacks [8, 41]; Ray’s Otherwise, they may become vulnerable to small subgroup renegotiation attack [45] is also an instance of this pattern. attacks [see e.g. 7, 46] which have been exploited in previous If an attacker A can see application-level protocol messages TLS attacks [55, 38]. Barker et al. [11] define a procedure between C and S, it can tunnel these messages through its for checking public keys, but we found that many TLS own connection with S, thereby impersonating C at S. implementations do not implement it. We analyzed TLS clients This attack is possible in three scenarios. First, if the and servers to check whether they accept degenerate public client C uses the same application-level credentials on en- keys (with small orders) like 0, 1, and −1; these keys always crypted and unencrypted channels. Second, if C uses the lead to pms ∈ {0, 1, −1}. While 0 and 1 are rejected by same credentials on different servers, one of which could be most implementation (to mitigate [38]), we found that NSS, malicious. Third, if C fails to correctly validate the server SChannel, and JSSE do accept −1. On the web, we found identity and confuses a malicious server A with an honest that all web browsers and about 12% of DHE-enabled servers server S. In all these cases, the application-level protocol of the top 10,000 Alexa list also accept −1. Such clients and should guarantee that the credentials released by C to A cannot servers are vulnerable to our key synchronization attack, since be used by A at S. the pms can be forced to be the same on both connections A common pattern to enforce this guarantee is to crypto- (with high probability), even if these clients and servers only graphically bind the (inner) application authentication to the accept known primes and correctly sample their keys. (outer) underlying TLS channel [8, 6, 49]. This binding helps The elliptic curve version of DHE (ECDHE) allows servers only inasmuch as the inner protocol employs strong keys (pub- to offer arbitrary curves, and so theoretically suffers from the lic or secret) or a passphrase-based challenge-response scheme resistant to dictionary attacks. Conversely, bearer tokens cannot exchanges are typically used to provide both server and client be protected. In this section, we discuss four such binding authentication, and hence, they both offer several ciphersuites mechanisms, and show how to break their guarantees using that do not use server certificates at all. the synchronizing TLS proxy of §V. The second precondition is that the client and server should be willing to accept new mutual identities during renegotiation. VI-A T HE T RIPLE H ANDSHAKE ATTACK . Suppose the at- Accepting a change of client identity (or client authentication tacker A has an anonymous-client TLS connection to server S. on an anonymous session) is one of the purposes of renegoti- When A tries to access a user-protected resource, S triggers ation, but accepting a change of server may seem unusual. We a renegotiation to require A to authenticate as a valid user, experimentally tested a wide variety of TLS client applications, with a client certificate or some other credential (PSK, SRP, including mainstream browsers, popular HTTPS libraries such etc.). This pattern is enabled, for example, on the Apache web as CURL, serf, and neon, version control systems, VPN clients, server, when a client tries to access a protected directory. mail clients, etc. We found that a vast majority of them silently A wants to authenticate to S as C (without C’s credentials). accept a change of server identity during renegotiation, and More generally, even if A has previously authenticated to S, thus are vulnerable to our impersonation attack. it wants to change its authenticated identity to C. Why does this not contradict proofs of the TLS handshake? Before explaining our attack, it is useful to recall the Most proofs [e.g. 35, 32] ignore renegotiation and resumption; 2009 renegotiation attack [45] and countermeasure [49], which [14] supports resumption but not renegotiation; [29] considers cryptographically binds each handshake on a connection to renegotiation but not resumption; [15] supports both but relies the preceding one, by passing the verify data of the previous on the application to correctly handle epoch changes. handshake (if there was one) in the client and server hellos of the new handshake. Therefore, if A initiates a full handshake Web Exploit and Mitigation As a concrete example, we with S, but later tries to forward C’s handshake to S as a implemented the above attack as a web server acting as a renegotiation, the verify data in C’s hello would not match A’s synchronizing proxy between a browser C and an honest handshake, prompting the server to reject the renegotiation. website S. After proxying the initial handshake and session What if a session is resumed on a new connection? The resumption, A can tamper with the connection in many ways, first handshake now is an abbreviated handshake; it only before instigating renegotiation: authenticates the session master secret, not the whole session. • A can send a POST message to S which will get Thus, the renegotiation countermeasure does nothing to bind subsequently attributed to C after renegotiation. the new connection to the old session. This re-enables the man- • A can send a page with JavaScript to C, so that the script in-the-middle impersonation attack it was meant to fix. gets executed later, in the client-authenticated session. Assume the adversary A has set up synchronized sessions • A can source a client-authenticated page from S in one and connections with C and S. If C resumes the session on a frame at C while reading its contents from another frame new connection, A can resume the same session on a new sourced at A, bypassing the same origin policy (XSS). connection to S. As discussed in §V-C, at the end of the All of these attacks can be used to subvert both user authentica- abbreviated handshake, the verify data on both connections is tion on the server and same-origin protections on the browser. the same. Now, if C or S initiates a client-authenticated TLS Protections like CSRF tokens and Content Security Policy do renegotiation, A can simply forward all messages from C to S not help since the page’s origin is no longer reliable. and back, making no changes. The client and server hellos will We have disclosed this vulnerability to a number of browser refer to the verify data from the abbreviated handshake and vendors. The easiest mitigation is for web browsers to refuse thus be accepted by both parties. This triple handshake across a change of server identity during renegotiation (since their two connections is depicted in Figure 3. UI can hardly convey a HTTPS mashup of several origins); At the end of the renegotiation, from TLS’s viewpoint, C some of them have already made this change in response to and S share a new mutually-authenticated session. A does not our report. For web servers and other HTTPS applications, have the keys to this new session, but it may have injected we believe that restricting peer certificate changes would be a data in both directions before the renegotiation, and this data good default as well, with a careful review of the UI and API may now be mistakenly attributed by C to S, and vice versa. design in the cases when the identity is expected to change. In other words, the TLS peer on the connection has changed, VI-B B REAKING C OMPOUND AUTHENTICATION IN T UN - and the application may not realize it, defeating the purpose NELED P ROTOCOLS . Wireless authentication protocols such of the secure renegotiation extension. as EAP-TLS [51], PEAP [42] and EAP-TTLS [27] are par- Preconditions and Variations The attack above works re- ticularly susceptible to man-in-the-middle attacks even over gardless of whether the renegotiation uses client certificates, TLS [8] because of the ease with which other wireless PSK, or SRP to authenticate the client, and even if the initial devices and rogue access points can fool naive clients into handshake also used client authentication. connecting to them [19]. To protect against such attacks, The main precondition is that the client be willing to use the some of these protocols adopted new compound authentication same authentication credentials on A and S. This is reasonable mechanisms [44] that cryptographically bind the inner EAP for public-key certificates, which are often used as universal authentication protocol with the outer TLS tunnel. identity assertions when issued by trusted CAs. For SRP or In PEAP, when the inner protocol is MSChapv2 [1] for PSK credentials, this may not seem as likely, but these key example, the inner protocol generates a session key (ISK) that is combined with a tunnel key (TK) generated from the The intent is that tls-unique be a unique representative of outer TLS connection’s master secret (and client and server the current epoch, shared only between the two peers who randoms) to derive a compound authentication key (CMK) and established the epoch. Our synchronized session resumption encryption key (CSK) for subsequent use between the wireless breaks it by establishing different connections with honest device and access point. The idea is that these keys will only peers that have the same tls-unique value. be known to devices that participated both in the outer TLS To see how this can be concretely exploited, consider the handshake and the inner EAP authentication. SCRAM-SHA-1-PLUS protocol [39] used in the SASL and GSS-API families of authentication mechanisms in a variety TK = prf(ms, “client EAP encryption”, cr|sr) of applications like messaging (XMPP), mail (SMTP, IMAP), CMK|CSK = prf (TK, ISK) and directory services (LDAP). SCRAM is a challenge- response protocol where the client and server store different PEAP also features fast reconnect, an API for TLS session keys (CKp , SKp ) derived from a user’s password (p), and use resumption: as it moves from one wireless access point to them to authenticate one another. When used over TLS, another and needs to reconnect, the client simply resumes its the first two messages contain client and server nonces and TLS session and skips the inner authentication protocol. In the tls-unique value for the underlying TLS connection. this case, ISK is set to 0s so the compound authentication The last two messages contain MACs over these values, for and encryption keys depend only on TK. This mechanism authentication and channel binding: presumes that the tunnel key is unique on every connection; our synchronizing TLS proxy breaks this assumption and leads 1. C→S: u, cn, tls-unique to a new attack. 2. S→C: cn, sn, s, i As usual, A sets up synchronized connections with C and S 3. C→S: cn, sn, ClientProof(CKp , log1,2,3 ) and forwards the untampered MSChapv2 exchange to let C 4. C→S: cn, sn, ServerSignature(SKp , log1,2,3 ) authenticate to S, negotiate ISK, combine it with TK, and derive CMK and CSK. Since A only knows TK, he cannot read or tamper In our attack, C establishes, then resumes a session with A, with any messages after the authentication. Nonetheless, if A who synchronizes a connection with S to have the same uses fast reconnect to resume the TLS session with S, the inner tls-unique value. A then forwards the SCRAM messages EAP authentication is skipped, and the new compound keys between C and S. Since the server identity is not part of are only derived from TK. Yet, S still associates the connection the exchange and the tls-unique values match, the SCRAM with C, resulting in a complete impersonation by A, without authentication succeeds, enabling A to impersonate C at S. any involvement from C. A precondition for the attack is that C be willing to accept Preconditions and Mitigations To make the attack work, A’s certificate, and this is already considered a security risk the malicious access point must convince the user to trust its for SCRAM-like protocols, since they then become vulnerable certificate, which can be achieved in a number of cases [19]. to dictionary attacks. However, the tls-unique protection is The mitigation for tunneled protocols is not straightforward. meant to protect users from impersonation even if the TLS At the TLS level, a more general mitigation would be to protocol uses an anonymous key exchange [39, §9]. Our attack change the master secret computation, as we discuss in §VII. In shows that this is not the case. PEAP, one possibility is to change the tunnel key computation To prevent this attack without impacting TLS, we recom- to include the server’s identity, represented by the server’s mend significant changes to the specification of tls-unique certificate or its hash: in §VII. With such modifications, tls-unique may possibly become truly unique across connections. TK = prf(ms, “client EAP encryption”, cr|sr|cert S ) VI-D B REAKING C HANNEL -B OUND T OKENS ON THE VI-C B REAKING TLS C HANNEL B INDINGS . Channel bind- W EB . Channel ID is a TLS extension [10], implemented by ings [56] are a generic protocol composition mechanism, Chrome and all Google servers, that aims to bind web au- whereby a transport-level cryptographic protocol such as IPsec, thentication tokens such as cookies to a cryptographic channel SSH, or TLS can expose specific session and connection between a client and a server, without the need for client parameters to applications, most notably to bind authentication certificates. A channel can be long-lived (at least as long as mechanisms to the underlying secure channel. Their stated cookies) and consists of many TLS sessions and connections. goal is to establish that “no man-in-the-middle exists between Channel ID is a follow-up to the previously published origin- two end-points that have been authenticated at one network bound certificates proposal of Dietz et al. [24], which was layer but are using a secure channel at a lower network considered impractical to implement and deploy. layer”. TLS implementations expose three channel bindings to A TLS client that supports Channel ID generates and applications [6]; we consider one of them here and another stores a public-private elliptic curve key pair (pk cid,S , sk cid,S ) (tls-server-end-point) in the online material. The associated to each domain name S that it connects to. The TLS ‘tls-unique’ channel binding for a given TLS connection handshake is modified so that, instead of a client certificate is defined as the first finished message in the most recent and certificate verify message, the client sends a Channel ID handshake on the connection. If the most recent handshake authentication message that contains the public key (a point is a full handshake, this value is the client verify data cvd ; if on the P-256 elliptic curve) and an ECDSA signature of the it is an abbreviated handshake, it is the server verify data svd . handshake log using the private key. To protect the privacy of the client’s public key from passive eavesdroppers, the the ciphersuite. For SSL3 and earlier versions of TLS, this is authentication message is sent encrypted after the client’s CCS the concatenation of MD5 and SHA1 hashes. We require that message, but this does not affect its authentication properties. TLS implementations compute and store tls-session-hash The main protocol goal is that, unlike bearer tokens, the within its session structure and expose it to implementations. client’s Channel ID cannot be used by a malicious server A Why this definition? We only hash messages up to the to impersonate the client on a different server S, even if C client key exchange, because at this point the negotiation is accidentally connects to A using its Channel ID for S. In complete and all the inputs to the master secret are available, fact, this should be impossible even if A obtains the private so most TLS implementations will create (but not cache) the key of a certificate valid for S, provided Channel ID is only session structure. Notably, the hashed log includes the nonces, enabled with forward-secret ciphersuites such as DHE [10, the ciphersuite, key exchange messages, client and server §6]. Consequently, an application that binds its tokens to the certificates, and any identities passed in protocol extensions. Channel ID make them unusable on a different TLS client Our definition of the hash functions matches those used without the associated private key. A typical example is for S for the finished messages in SSL3 and TLS 1.0–1.2; hence, to create a cookie by signing the session identifier with the implementations already keep a running hash of the log and Channel ID public key: we just re-use its value. Implementing this channel binding c = signed(sk S , [sid , pk cid ]) increases the cached session size by a single hash, and has no performance impact. S would then only accept this cookie over a TLS connection We define a new hash value instead of reusing the client or authenticated by sk cid , so stealing the cookie is of no use. server verify data for three reasons. (1) It is compatible with Attack and Mitigation The security of Channel ID relies on stateless servers [50], which must send the session ticket before the uniqueness of the handshake log (log c ). If the attacker A the server finished message, so the server verify data is not can create a session to S with the same log, it can reuse C’s available yet. (2) Being longer than the verify data, the session Channel ID signature to impersonate C at S. Our synchroniz- hash offers stronger collision resistance. While collisions may ing proxy achieves exactly this feat after resumption. be less problematic for (the usually few) renegotiations on a Suppose C establishes, then resumes a TLS session with A. single connection, a session can be long-lived and frequently A can synchronize a connection to S such that the log resumed. (3) We could have reused the input to the client in the resumption handshake is identical between C-A and verify data, but it would not offer any clear advantages, and our A-S. Hence, the Channel ID signature on the resumption current definition is more suitable for our proposed extensions. handshake can be replayed to S, allowing A to successfully Recommended Usage We recommend that protocols such as impersonate C. Henceforth, A can obtain S’s channel-bound SCRAM use tls-session-hash rather than tls-unique cookies meant for C and freely use them on this connection. for channel binding. To fix Channel ID, we recommend This attack is well within the threat model of Channel ID. The that the signature on abbreviated handshakes include the Channel ID authors promptly responded to our report and in tls-session-hash of the resumed session. To derive ap- response, the protocol specification is being revised to include plication keys from the master secret, like in PEAP, we the hash of the original handshake in the Channel ID signature recommend adding tls-session-hash to the PRF. of abbreviated handshakes. VII-B C ONTEXT B INDING FOR M ASTER S ECRETS . We pro- VII. C OUNTERMEASURES pose a new extension for all versions of TLS and DTLS that causes negotiated session parameters to be included in the We propose several countermeasures at the TLS level that master secret computation, following the principle of context prevent our man-in-the-middle attacks at their source with binding [21], whereby computed keys should be usage-specific. few or no changes required to application-level mechanisms. The ideas behind these proposals emerged from discussions As usual, the extension is signaled in the client and server with various implementors and protocol experts and we are hello messages; if both peers support it, the handshake pro- cautiously optimistic about their adoption. Since new protocol ceeds as usual, except that the master secret is computed as: extensions and features can take a long time to propagate, we ms = prf(pms, “extended master secret”, also discuss short-term mitigations for various applications. tls-session-hash) VII-A A NEW CHANNEL BINDING . In §V-C and §VI-C, we found that neither the session id, nor the master secret, nor the The inclusion of tls-session-hash, instead of just the tls-unique channel binding served as unique representatives pair of nonces, ensures that the resulting master secret depends for a TLS session. Hence, we propose a new TLS channel on all the negotiated session parameters. The master secret im- binding, called tls-session-hash, that captures all the plicitly authenticates these parameters, and different sessions negotiated parameters for a session. will have different master secrets, foiling our attacks. We define tls-session-hash for a given TLS session as We find this solution elegant since it protects all TLS the hash of the handshake messages up to and including the handshake modes (including RSA and DHE) and protocol client key exchange message in the original handshake that versions, and it allows application-level protocols like PEAP created the session. The hash function used depends on the to safely use the TLS master-secret without any changes. The protocol version. For TLS 1.2, this is the hash function in idea of including additional materials in the master secret computation is not new [21, 31, 3] but our proposal merits (We refer to the online materials for a more detailed de- more detailed analysis, which we leave for future work. scription of miHTTPS, its code, and its verification.) VII-C S ECURE R ESUMPTION I NDICATION . We propose a Secure Channels Our main communication abstraction is a mandatory extension for all versions of TLS and DTLS that long-term, stateful channel between a client and a host. Each complements the renegotiation indication extension [49] by client may create any number of channels and use them to also protecting session resumption across multiple connections. request documents from URLs at different hosts; each channel As in [49], the extension is signaled in the client and supports parallel requests, as required e.g. when loading a server hello messages (see §IV-C), but only when propos- web page that includes numerous resources. Each request may ing and accepting resumption, respectively. It contains the asynchronously return a document (in any order). tls-session-hash value of the session being resumed. Such channels are not reliable: requests and responses may Peers supporting the extension must check that this value get lost or delayed, and their sender have no explicit acknowl- matches the one recorded in their locally stored session before edgment of peer reception. Instead, responses confirm requests, proceeding with the abbreviated handshake. The exchanged and cookies attached to requests confirm prior responses. session hashes are authenticated by the master secret in the In the command line, the host=NAME option indicates finished messages of the resumption, cryptographically binding that a new channel should be created and its ID returned, the new connection to the resumed session. If one of the peers whereas channel=ID indicates the local identifier of an does not support the extension, the other should refuse session existing channel to reuse. These application-level channels are resumption and may instead offer a full handshake. not primitive in HTTPS or TLS; they intuitively account for VII-D S UMMARY OF M ITIGATIONS . We implemented the a series of related requests issued by a client. For example, a session hash channel binding and our two extensions as patches user may have long-lived authenticated channels to every host to OpenSSL and miTLS, and we tested their interoperability she trusts, plus shorter-lived anonymous channels. The server for all versions of TLS and DTLS. Our patches fit well into is always authenticated. The user may use the client=NAME the code structure and have no visible effect on performance. option, where NAME refers to a valid client certificate she owns Independently, applications that rely on existing TLS APIs to be used to authenticate her requests on the channel. can mitigate the attacks of this paper by following some con- Simplifications We associate a unique host name to each servative design principles, at some cost to their functionality. channel, treating each host as a separate principal: thus, we 1) Do not allow the peer to renegotiate its certificate. do not deal with related sub-domains, redirects, or wildcards 2) Do not use tls-unique after session resumption . in certificate names. We also do not support mixtures of HTTP 3) To derive application keys from the TLS master secret, and HTTPS. Thus, we avoid many complications with cookies hash the session’s certificates into the derivation. discussed in §II and §III. (Applications may still multiplex 4) Buffer application data until its semantics is unambigu- between hosts and protocols on top of our interface—what ous; discard it if the TLS connection is torn down. matters is that we do not share private state between channels.) 5) Do not share secret cookies between HTTP and HTTPS Client and Server Credentials We rely on the public-key connections, or between different origins. infrastructure for X.509 certificates, and require that client and host names exactly match their certificates’ common names. VIII. V ERIFIED A PPLICATION S ECURITY OVER TLS Our threat model does not cover certificates mis-issued to the VIII-A MI HTTPS: A BASIC HTTPS C LIENT . To validate adversary, or issued for different purposes with a common our application-level recommendations and show that one can name that matches an honest principal. indeed achieve transparent application-level security over TLS, Credentials are associated with the whole channel, once we build and verify an exemplary HTTPS library, at the same and for all. The host name cannot be changed, preventing level of abstraction as the CURL library, for example, but with the renegotiation attack of §VI-A. The client can decide to fewer features. Its client command-line interface is as follows: authenticate later on an anonymous channel, and from the $ mihttps --help server’s viewpoint, this suffices to attribute all requests on the Usage: mihttps [options] REQUEST channel to that client. From the client’s viewpoint, binding her --host=NAME https server host name name to the channel before a particular request guarantees that --channel=ID channel identifier the server will only process it after client authentication. --client=NAME authenticated client name Local State and Cookies Our channels maintain local, private Our goal is to provide (1) a basic API with strong implicit state, including e.g. open connections, live sessions, cookies, security; and (2) a flexible implementation that supports typical and the names associated with the channel. Our channels also mechanisms available in HTTP (cookies) and TLS (multiple buffer request and response fragments, in order to deliver only connections, renegotiation, resumption, late client authentica- whole HTTPS messages to the application—this simply foils tion). miHTTPS consists of 600 lines of F# coded on top of the truncation attacks, including those of §III-B. miTLS verified reference implementation [15]. In particular, At the server, we partition incoming requests into sepa- our client automatically processes HTTP 1.0 headers, cookies, rate channels and track requests received from each client etc, and interoperates with existing, unmodified web servers. by attaching a (locally stored) fresh random cookie to each We tested e.g. authenticated webmail access to Roundcube. response. The set of responses actually received can then be inferred from the cookies attached to latter requests. (Assum- 1 type name = string (∗ common names for both clients & hosts ∗) ing sufficient cookie storage space and entropy to prevent 2 type (;host:name) chan 3 predicate Honest of name (∗ no compromised certificate ∗) collisions, this pattern provides accurate tracking information.) 4 predicate Client of name ∗ host:name ∗ (;host)chan VIII-B S ECURITY G OALS (I NFORMAL ). We primarily focus 5 6 module Data (∗ defined by the application ∗) on application-level channel integrity—see the online version 7 type (;host,chan)request for privacy. We follow the cryptographic model of [15] and 8 type (;host,chan,request)document configure honest clients and servers to only negotiate strong 9 module Certificate (...) ciphersuites and algorithms [as defined by 15]. We show that, 10 module Server (...) with overwhelming probability, the following properties hold: 11 module Client 12 val create: h:name → (;h) chan 1) Request Integrity: when an honest server accepts a 13 val request: h:name → c:((;h)chan) → request and attributes it to a channel bound to honest 14 (a:name{Client(c,a)})option → r:(;h,c)request → unit server and client names, the client has indeed sent the 15 val poll: h:name → c:((;h)chan) → 16 (r:(;h,c)request ∗ (;h,c,r)document) option request on that channel, with matching principal names. 2) Response Integrity: when an honest client accepts a Fig. 4. miHTTPS interface (excerpt) document in reply to a request to an honest server, that server has indeed sent the document in response to this specify the host, the channel, and the request (for responses), request. (This property is sometimes called correlation.) so only the application above miHTTPS can create and access 3) Tracking: when an honest server accepts a request values at those types. They yield strong, information-theoretic echoing the cookie of a response on a channel with an security: provided that the channel is between honest client and honest client, the client indeed received this response. server, type safety ensures that our protocol stack, including Property 1 excludes any mis-attribution of a request to a HTTPS, TLS, TCP, and any network adversary, cannot read client. Properties 1 and 2 apply to whole messages, thereby ex- their content (except for their size after encryption), tamper cluding truncations. This is achieved by parsing and buffering with their content, or move contents from one channel to message fragments until the whole message has been received, another. Essentially, the protocol can only pass requests un- decrypted, and authenticated. changed from clients to servers, and similarly for responses. The Certificate module manages certificates. Reflecting our VIII-C MI HTTPS: S ECURE T YPED I NTERFACE . We follow threat model, it has functions for generating certificates for the modular type-based cryptographic verification method [26] Honest names and endorsing keys for dishonest names. that was used to obtain the main security theorem for the The Server module defines the API for miHTTPS servers. miTLS API [15]. They specify computational security for var- The Client module is the actual API used by client appli- ious constructions and protocols using precise typed interfaces cations, such as our command-line client. It has functions for (instead of code-based games or ideal functionalities). They creating a new channel towards a fixed host h, for sending employ an expressive refinement-based type system for F#, requests (with optional client authentication), and for polling write detailed typed annotations (4,000 lines for miTLS), and responses to prior requests. These functions have precise value- verify their code against them automatically using F7, an dependent types specifying their pre- and post-conditions. For extended typechecker, coupled with Z3, an SMT solver. instance, request takes 4 parameters: the target host h; an The verification effort for miHTTPS consists of specifying existing channel c for that host; an optional client name a its typed API and letting F7 typecheck its 600 lines of code, authorized by the user for that channel (as indicated by the using the lower-level, verified, precisely-typed API of miTLS. predicate Client(c,a)); and a request for that host and channel. In the rest of the section, we outline the types we use to capture the security goals of §VIII-B. Figure 4 shows fragments of our typed specification for IX. I MPACT AND L IMITATIONS miHTTPS, focusing on the main functions for the client. It We have presented a series of attacks on authentication defines a type for names—plain strings used as common names mechanisms built within and over TLS. Table II summarizes in certificates—and for channels: type (;host:name)chan. This these new attacks and compares them to previous attacks, in type is indexed by a value, host, itself of type name, recording terms of their impact and limitations. The table lists precondi- in the type that the channel should be used only for com- tions for each attack: what the attacker must be capable of; munications with servers with a valid certificate for host. This how the application (mis-)uses TLS; and whether previous type is also abstract, hiding its representation, so that only our mitigations block the attack () or not (). miHTTPS implementation can access it; applications can just For example, the second row indicates that the cookie cutter pass channels as arguments to the API, but they cannot access attack of §III-B requires a network attacker and a client their internal states (and so cannot accidentally leak keys) or application that processes truncated HTTP headers over TLS modify the host index (and so cannot get confused between and a server application that allows chosen plaintexts before channels to different hosts). the Set-Cookie header. Its advantage over previous TLS Our API has 3 main modules, and is parameterized by an truncation attacks is a higher impact: it enables full HTTPS application module, Data, provided by the application, that session hijacking (by stealing session cookies) between main- defines types for requests (URLs) and responses (documents). stream web browsers and popular websites such as Google and These types are both abstract and indexed. Their indexes Facebook. Conversely, our variant of network-based session TABLE II. S UMMARY OF ATTACKS : NOVELTY, IMPACT AND PRECONDITIONS Attacker Abilities API Assumptions Mitigations Attack Broken Mechanism 1 2 3 4 5 6 7 8 9 10 11 Refs TLS Truncation HTTPS Session (Tampered) [13, 52] ∗ Cookie Cutter HTTPS Session (Hijacked) §III-B Session Forcing (Server) HTTPS Session (Login CSRF) [12, 18] Session Forcing (Net) ∗ Truncation+Session Forcing HTTPS Session (Login CSRF) §III-C TLS Renegotiation (Ray) TLS Client Auth (Certificate) [49, 45] TLS Renegotiation (Rex) ∗ Triple Handshake (RSA) §VI-A ∗ TLS Client Auth (Certificate) Triple Handshake (DHE) §V-B MITM Tunnel Auth (Net) EAP (Certificate, Password) [8] MITM Tunnel Auth (Server) EAP (Certificate) ∗ MITM Compound Auth EAP (Certificate) §VI-B ∗ MITM Channel Bindings SASL (SCRAM-Password) §VI-C ∗ MITM Channel ID Channel ID (Public-Key) §VI-D 1. Client connects to untrusted server 17. Client accepts unknown DH groups/degenerate public keys 2. Active network attacker 18. Client accepts server certificate change during renegotitation 3. Client authenticates on untrusted server 19. HSTS: Require TLS for all actions on trusted server 4. Attacker controls one subdomain on trusted server 10. Require renegotiation indication extension 5. Application accepts truncated TLS streams 11. Bind authentication protocol to TLS channel 6. Application sends attacker-chosen plaintext in channel forcing (fifth row, §III-C) has the same impact as previous certificate change during renegotiation prevented in attacks; its novelty is that it bypasses their HSTS mitigation. CVE-2013-6628. Our new attacks on TLS renegotiation, PEAP, SASL, and • SChannel (used by Internet Explorer): Degenerate Channel ID are server-based man-in-the-middle attacks. They Diffie-Hellman public keys and server certificate change require that a client be willing to connect and authenticate during renegotiation both prevented by a security update. with some credential (e.g. an X.509 certificate) at an untrusted • NSS (used by Firefox): Degenerate Diffie-Hellman pub- server. The resulting attack is that the untrusted server can lic keys prevented in CVE-2014-1491. impersonate the client at any trusted server that accepts the • Channel ID (implemented in Chrome): Impersonation same credential. The precondition that the client be willing to attack prevented by using only ECDHE ciphersuites; use its credential at an untrusted server is restrictive: it is more specification revised to use session hashes (§VII-A). reasonable for public-key certificates than for server-specific • Safari: Notified of header truncation attack on June 13, tokens such as passwords. Still, such man-in-the-middle at- 2013. Notified of an incorrect renegotiation behavior on tacks by malicious servers were meant to be prevented by January 10, 2014, which was fixed in a later update. various channel-binding mechanisms built into these protocols, • Apache: Notified of POST message truncation in and our attacks show that these mitigations are insufficient. mod_php on April 29, 2013. Acknowledged, not fixed. Our triple handshake attack on TLS renegotiation (§VI-A) These short-term fixes, however, do not address our attacks bypasses the renegotiation indication countermeasure, but it on channel bindings in SASL and compound authentication applies only to servers that authenticate clients with certifi- in PEAP. More generally, our findings falsify the assumptions cates during renegotiation. Such server configurations are not made by the authors and users of various protocol specifi- widespread, but can still be found in banks, certificate authori- cations [23, 49, 48, 51, 27, 42, 1, 6, 39, 33, 10]. A more ties, and VPN services. Furthermore, our impersonation attacks systematic fix would be to strengthen the TLS protocol itself apply only to clients that are willing to accept a change of to provide these stronger expected authentication properties. server certificates during renegotiation. Our experiments show We contacted various members of the TLS working group, that these and other preconditions in the table are frequently including authors of the renegotiation extension [49]. They met by popular web browsers and TLS and HTTPS libraries. acknowledged the attack and we are collaborating on two internet drafts that describe the mechanisms proposed in §VII. IX-A R ESPONSIBLE D ISCLOSURE . We reported the attacks We informed authors of TLS channel bindings [6] of our to several software vendors and suggested short-term fixes that attacks and they acknowledged that tls-unique in its current invalidate the preconditions of these attacks. We summarize form should not be used after resumption. Discussions on their responses below. In light of our findings, we advocate revising the channel binding specification are ongoing. that all applications that rely on TLS carefully review their The security of our proposed extensions remains to be use of TLS libraries and implement similar fixes if necessary. formally evaluated. We plan to extend the cryptographic proofs • Chromium (used by Chrome, Android, Opera): Header of miTLS to precisely model these extensions and verify that truncation attacks prevented in CVE-2013-2853. Server they provide stronger authentication guarantees for TLS. ACKNOWLEDGEMENTS [28] M. Georgiev, S. Iyengar, S. Jana, R. Anubhai, D. Boneh, and V. Shmatikov. The most dangerous code in the world: validating SSL We thank Martı́n Abadi, Bruno Blanchet, Catalin Hritcu, certificates in non-browser software. In ACM CCS, 2012. Markulf Kohlweiss, Adam Langley, Marsh Ray, Martin Rex, [29] F. Giesen, F. 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