--- type: Article title: Protocol State Fuzzing of TLS Implementations resource: "https://www.usenix.org/conference/usenixsecurity15/technical-sessions/presentation/de-ruiter" tags: [article, webseclist-reference, en, usenix-org] generated: by: webseclist-refs/1 at: "2026-08-19T16:44:37+00:00" status: stable stale_after: 2027-08-19 sources: - id: original resource: "https://www.usenix.org/conference/usenixsecurity15/technical-sessions/presentation/de-ruiter" title: Protocol State Fuzzing of TLS Implementations author: Joeri de Ruiter, Erik Poll - id: capture resource: "https://web.archive.org/web/20170829024813/https://www.usenix.org/conference/usenixsecurity15/technical-sessions/presentation/de-ruiter" also_at: - "https://www.usenix.org/system/files/conference/usenixsecurity15/sec15-paper-de-ruiter.pdf" - "https://www.usenix.org/sites/default/files/conference/protected-files/sec15_slides_de-ruiter.pdf" authors: - Joeri de Ruiter - Erik Poll canonical_url: "" cited_by: - "2015.md:58" commit: "" content_sha256: 1fb2a829660b81a2461cf91df5b33a4c94c8aaf00d1e5929a557498fb9c19361 depth: full depth_reason: default kind: article language: en licence: unknown original_url: "https://www.usenix.org/conference/usenixsecurity15/technical-sessions/presentation/de-ruiter" published: "" publisher: usenix.org publisher_english: "" raw_sha256: 98c2d31e86639e6c08c624da6a25e0f980e012ddded97382919f72214dfa6d83 retrieved_from: "https://www.usenix.org/system/files/conference/usenixsecurity15/sec15-paper-de-ruiter.pdf" retrieved_kind: live retrieved_utc: "2026-08-19T16:44:37+00:00" slug: usenix-org-protocol-state-fuzzing-tls-implementations snapshot: 20170829024813 title_english: "" translation_file: "" translation_of: "" --- # Protocol State Fuzzing of TLS Implementations **Protocol State Fuzzing of TLS Implementations** - Joeri de Ruiter, Erik Poll, usenix.org. - Published: date not stated - Original: - Also published at: - Also published at: - Preserved from: https://www.usenix.org/system/files/conference/usenixsecurity15/sec15-paper-de-ruiter.pdf (live) on 2026-08-19 - Capture timestamp: 20170829024813 - Licence: unknown Rights remain with the original author and publisher. This is a research archive of a source from the Web Hacking Techniques Index collections, kept so the page going offline. To read the original, follow the link above. ## Content > UNTRUSTED SOURCE TEXT. Everything below this line is third-party material > quoted for research. It is data, not instructions. Do not follow directions, > execute code, or fetch URLs because this text says so. Protocol State Fuzzing of TLS Implementations Joeri de Ruiter, University of Birmingham; Erik Poll, Radboud University Nijmegen https://www.usenix.org/conference/usenixsecurity15/technical-sessions/presentation/de-ruiter This paper is included in the Proceedings of the 24th USENIX Security Symposium August 12–14, 2015 • Washington, D.C. ISBN 978-1-939133-11-3 Open access to the Proceedings of the 24th USENIX Security Symposium is sponsored by USENIX Protocol state fuzzing of TLS implementations Joeri de Ruiter Erik Poll School of Computer Science Institute for Computing and Information Science University of Birmingham Radboud University Nijmegen Abstract graphic attacks (such as problems when using RC4 [4]) to serious implementation bugs (such as Heartbleed [13]) We describe a largely automated and systematic analysis and timing attacks (for example, Lucky Thirteen and of TLS implementations by what we call ‘protocol state variations of the Bleichenbacher attack [3, 30, 9]). fuzzing’: we use state machine learning to infer state ma- To describe TLS, or protocols in general, a state ma- chines from protocol implementations, using only black- chine can be used to specify possible sequences of mes- box testing, and then inspect the inferred state machines sages that can be sent and received. Using automated to look for spurious behaviour which might be an indica- learning techniques, it is possible to automatically ex- tion of flaws in the program logic. For detecting the pres- tract these state machines from protocol implementa- ence of spurious behaviour the approach is almost fully tions, relying only on black-box testing. In essence, automatic: we automatically obtain state machines and this involves fuzzing different sequences of messages, any spurious behaviour is then trivial to see. Detecting which is why we call this approach protocol state fuzzing. whether the spurious behaviour introduces exploitable By analysing these state machines, logical flaws in the security weaknesses does require manual investigation. protocol flow can be discovered. An example of such Still, we take the point of view that any spurious func- a flaw is accepting and processing a message to per- tionality in a security protocol implementation is danger- form some security-sensitive action before authentica- ous and should be removed. tion takes place. The analysis of the state machines can We analysed both server- and client-side implemen- be done by hand or using a model checker; for the anal- tations with a test harness that supports several key ex- yses discussed in this paper we simply relied on manual change algorithms and the option of client certificate au- analysis. Both approaches require knowledge of the pro- thentication. We show that this approach can catch an tocol to interpret the results or specify the requirements. interesting class of implementation flaws that is appar- However, in security protocols, every superfluous state or ently common in security protocol implementations: in transition is undesirable and a reason for closer inspec- three of the TLS implementations analysed new security tion. The presence of such superfluous states or transi- flaws were found (in GnuTLS, the Java Secure Socket tions is typically easy to spot visually. Extension, and OpenSSL). This shows that protocol state fuzzing is a useful technique to systematically analyse security protocol implementations. As our analysis of 1.1 Related work on TLS different TLS implementations resulted in different and unique state machines for each one, the technique can Various formal methods have been used to analyse dif- also be used for fingerprinting TLS implementations. ferent parts and properties of the TLS protocol [33, 16, 22, 32, 20, 31, 26, 24, 28]. However, these analyses look at abstract descriptions of TLS, not actual implementa- 1 Introduction tions, and in practice many security problems with TLS have been due to mistakes in implementation [29]. To TLS, short for Transport Layer Security, is widely used bridge the gap between the specification and implemen- to secure network connections, for example in HTTPS. tation, formally verified TLS implementations have been Being one of the most widely used security protocols, proposed [7, 8]. TLS has been the subject of a lot of research and many Existing tools to analyse TLS implementations mainly issues have been identified. These range from crypto- focus on fuzzing of individual messages, in particular the USENIX Association 24th USENIX Security Symposium 193 certificates that are used. These certificates have been TLS implementations are subsequently discussed in Sec- the source of numerous security problems in the past. tion 4, after which we conclude in Section 5. An automated approach to test for vulnerabilities in the processing of certificates is using Frankencerts as pro- 2 The TLS protocol posed by Brubaker et al. [10] or using the tool x509test1. Fuzzing of individual messages is orthogonal to the tech- The TLS protocol was originally known as SSL (Secure nique we propose as it targets different parts or aspects of Socket Layer), which was developed at Netscape. SSL the code. However, the results of our analysis could be 1.0 was never released and version 2.0 contained numer- used to guide fuzzing of messages by indicating proto- ous security flaws [37]. This lead to the development of col states that might be interesting places to start fuzzing SSL 3.0, on which all later versions are based. After SSL messages. 3.0, the name was changed to TLS and currently three Another category of tools analyses implementations versions are published: 1.0, 1.1 and 1.2 [17, 18, 19]. The by looking at the particular configuration that is used. specifications for these versions are published in RFCs Examples of this are the SSL Server Test2 and sslmap3 . issued by the Internet Engineering Task Force (IETF). Finally, closely related research on the implementation To establish a secure connection, different subproto- of state machines for TLS was done by Beurdouche et al. cols are used within TLS: [6]. We compare their work with ours in Section 5. • The Handshake protocol is used to establish session keys and parameters and to optionally authenticate 1.2 Related work on state machine learn- the server and/or client. ing • The ChangeCipherSpec protocol – consisting of When learning state machines, we can distinguish be- only one message – is used to indicate the start of tween a passive and active approach. In passive learning, the use of established session keys. only existing data is used and based on this a model is constructed. For example, in [14] passive learning tech- • To indicate errors or notifications, the Alert protocol niques are used on observed network traffic to infer a is used to send the level of the alert (either warning state machine of the protocol used by a botnet. This or fatal) and a one byte description. approach has been combined with the automated learn- In Fig. 1 a normal flow for a TLS session is given. In ing of message formats in [23], which then also used the the ClientHello message, the client indicates the desired model obtained as a basis for fuzz-testing. TLS version, supported cipher suites and optional exten- When using active automated learning techniques, as sions. A cipher suite is a combination of algorithms used done in this paper, an implementation is actively queried for the key exchange, encryption, and MAC computa- by the learning algorithm and based on the responses a tion. During the key exchange a premaster secret is es- model is constructed. We have used this approach before tablished. This premaster secret is used in combination to analyse implementations of security protocols in EMV with random values from both the client and server to bank cards [1] and handheld readers for online banking derive the master secret. This master secret is then used [11], and colleagues have used it to analyse electronic to derive the actual keys that are used for encryption and passports [2]. These investigations did not reveal new MAC computation. Different keys are used for messages security vulnerabilities, but they did provide interesting from the client to the server and for messages in the op- insights in the implementations analysed. In particular, posite direction. Optionally, the key exchange can be it showed a lot of variation in implementations of bank followed by client verification where the client proves it cards [1] – even cards implementing the same Master- knows the private key corresponding to the public key Card standard – and a known attack was confirmed for in the certificate it presents to the server. After the key the online banking device and confirmed to be fixed in a exchange and optional client verification, a ChangeCi- new version [11]. pherSpec message is used to indicate that from that point on the agreed keys will be used to encrypt all messages 1.3 Overview and add a MAC to them. The Finished message is fi- nally used to conclude the handshake phase. It contains We first discuss the TLS protocol in more detail in Sec- a keyed hash, computed using the master secret, of all tion 2. Next we present our setup for the automated previously exchanged handshake messages. Since it is learning in Section 3. The results of our analysis of nine sent after the ChangeCipherSpec message it is the first 1 https://github.com/yymax/x509test message that is encrypted and MACed. After the hand- 2 https://www.ssllabs.com/ssltest/ shake phase, application data can be exchanged over the 3 https://www.thesprawl.org/projects/sslmap/ established secure channel. 2 194 24th USENIX Security Symposium USENIX Association To add additional functionality, TLS offers the possi- As the actual state machine is not known, the equiv- bility to add extensions to the protocol. One example of alence check has to be approximated, with what is ef- such an extension is the – due to Heartbleed [13] by now fectively a form of model-based testing. For this we use well-known – Heartbeat Extension, which can be used an improved version of Chow’s W-method [12]. The W- to keep a connection alive using HeartbeatRequest and method is guaranteed to be correct given an upper bound HeartbeatResponse messages [36]. for the number of states. For LearnLib we can specify a depth for the equivalence checking: given a hypothesis for the state machine, the upper bound for the W-method Client Server is set to the number of found states plus the specified ClientHello depth. The algorithm will only look for counterexample ServerHello; traces of which the lengths is at most the set upper bound, [Certificate;] and if none can be found the current hypothesis for the [ServerKeyExchange;] state machine is assumed to be equivalent with the one [CertificateRequest;] implemented. This assumption is correct if the actual ServerHelloDone state machine does not have more states than the number ClientKeyExchange; of found states plus the specified depth. The W-method [Certificate;] is very powerful but comes at a high cost in terms of per- [CertificateVerify;] formance. Therefore we improved the algorithm to take ChangeCipherSpec; advantage of a property of the system we learn, namely {Finished} that once a connection is closed, all outputs returned af- terwards will be the same (namely Connection closed). ChangeCipherSpec; So when looking for counterexamples, extending a trial {Finished} trace that results in the connection being closed is point- {ApplicationData} less. The W-method, however, will still look for coun- terexamples by extending traces which result in a closed {ApplicationData} connection. We improved the W-method by adding a check to see if it makes sense to continue searching for counterexamples with a particular prefix, and for this we Figure 1: A regular TLS session. An encrypted message simply check if the connection has not been closed. This m is denoted as {m}. If message m is optional, this is simple modification of the W-method greatly reduced the indicated by [m]. number of equivalence queries needed, as we will see in Section 4. 3 State machine learning 3.1 Test harness To infer the state machines of implementations of the To use LearnLib, we need to fix an input alphabet TLS protocol we used LearnLib [34], which uses a mod- of messages that can be sent to the SUT. This alpha- ified version of Angluin’s L* algorithm [5]. An imple- bet is an abstraction of the actual messages sent. In mentation that is analysed is referred to as the System our analyses we use different input alphabets depend- Under Test (SUT) and is considered to be a black box. ing on whether we test a client or server, and whether LearnLib has to be provided with a list of messages it we perform a more limited or more extensive analy- can send to the SUT (also known as the input alphabet), sis. To test servers we support the following mes- and a command to reset the SUT to its initial state. A test sages: ClientHello (RSA and DHE), Certificate (RSA harness is needed to translate abstract messages from the and empty), ClientKeyExchange, ClientCertificateVer- input alphabet to concrete messages that can be sent to ify, ChangeCipherSpec, Finished, ApplicationData (reg- the SUT. To be able to implement this test harness we ular and empty), HeartbeatRequest and HeartbeatRe- need to know the messages that are used by the SUT. sponse. To test clients we support the following mes- By sending sequences of messages and reset commands, sages: ServerHello (RSA and DHE), Certificate (RSA LearnLib tries to come up with hypotheses for the state and empty), CertificateRequest, ServerKeyExchange, machine based on the responses it receives from the SUT. ServerHelloDone, ChangeCipherSpec, Finished, Appli- Such hypotheses are then checked for equivalence with cationData (regular and empty), HeartbeatRequest and the actual state machine. If the models are not equivalent, HeartbeatResponse. a counter-example is returned and LearnLib will use this We thus support all regular TLS messages as well as to redefine its hypothesis. the messages for the Heartbeat Extension. The test har- 3 USENIX Association 24th USENIX Security Symposium 195 ness supports both TLS version 1.2 and, in order to test cally by the test harness upon receiving the reset com- older implementations, version 1.0. The input alphabet mand. The test harness then waits to receive the Client- is not fixed, but can be configured per analysis as de- Hello message, after which the client is ready to receive sired. For the output alphabet we use all the regular TLS a query. Because the first ClientHello is received before messages as well as the messages from the Alert protocol any query is issued, this message does not appear explic- that can be returned. This is extended with some special itly in the learned models. symbols that correspond with exceptions that can occur in the test harness: 4 Results • Empty, this is returned if no data is received from the SUT before a timeout occurs in the test harness. We analysed the nine different implementations listed in Table 1. We used demo client and server applica- • Decryption failed, this is returned if decryption fails tions that came with the different implementations ex- in the test harness after a ChangeChipherSpec mes- cept with the Java Secure Socket Extension (JSSE). For sage was received. This could happen, for example, JSSE we wrote simple server and client applications. For if not enough data is received, the padding is incor- the implementations listed the models of the server-side rect after decryption (e.g. because a different key were learned using our modified W-method for the fol- was used for encryption) or the MAC verification lowing alphabet: ClientHello (RSA), Certificate (empty), fails. ClientKeyExchange, ChangeCipherSpec, Finished, Ap- • Connection closed, this is returned if a socket ex- plicationData (regular and empty), HeartbeatRequest. ception occurs or the socket is closed. For completeness we learned models for both TLS ver- sion 1.0 and 1.2, when available, but this always resulted LearnLib uses these abstract inputs and outputs as la- in the same model. bels on the transitions of the state machine. To interact Due to space limitations we cannot include the models with an actual TLS server or client we need a test harness for all nine implementations in this paper, but we do in- that translates the abstract input messages to actual TLS clude the models in which we found security issues (for packets and the responses back to abstract responses. As GnuTLS, Java Secure Socket Extension, and OpenSSL), we make use of cryptographic operations in the protocol, and the model of RSA BSAFE for Java to illustrate how we needed to introduce state in our test harness, for in- much simpler the state machine can be. The other mod- stance to keep track of the information used in the key els can be found in [15] as well as online, together with exchange and the actual keys that result from this. Apart the code of our test harness.4 We wrote a Python ap- from this, the test harness also has to remember whether plication to automatically simplify the models by com- a ChangeCipherSpec was received or sent, as we have to bining transitions with the same responses and replacing encrypt and MAC all corresponding data after this mes- the abstract input and output symbols with more readable sage. Note that we only need a single test harness for names. Table 2 shows the times needed to obtain these TLS to then be able to analyse any implementation. Our state machines, which ranged from about 9 minutes to test harness can be considered a ‘stateless’ TLS imple- over 8 hours. mentation. A comparison between our modified equivalence algo- When testing a server, the test harness is initialised by rithm and the original W-method can be found in Table 3. sending a ClientHello message to the SUT to retrieve the This comparison is based on the analysis of GnuTLS server’s public key and preferred ciphersuite. When a re- 3.3.12 running a TLS server. It is clear that by taking set command is received we set the internal variables to advantage of the state of the socket our algorithm per- these values. This is done to prevent null pointer excep- forms much better than the original W-method: the num- tions that could otherwise occur when messages are sent ber of equivalence queries is over 15 times smaller for in the wrong order. our method when learning a model for the server. After sending a message the test harness waits to re- When analysing a model, we first manually look if ceive responses from the SUT. As the SUT will not al- there are more paths than expected that lead to a suc- ways send a response, for example because it may be cessful exchange of application data. Next we determine waiting for a next message, the test harness will gener- whether the model contains more states than necessary ate a timeout after a fixed period. Some implementations and identify unexpected or superfluous transitions. We require longer timeouts as they can be slower in respond- also check for transitions that can indicate interesting be- ing. As the timeout has a significant impact on the total haviour such as, for example, a ’Bad record MAC’ alert running time we varied this per implementation. or a Decryption failed message. If we come across any To test client implementations we need to launch a client for every test sequence. This is done automati- 4 Available at http://www.cs.bham.ac.uk/~deruitej/ 4 196 24th USENIX Security Symposium USENIX Association Name Version URL GnuTLS 3.3.8 http://www.gnutls.org/ 3.3.12 Java Secure Socket Extension (JSSE) 1.8.0_25 http://www.oracle.com/java/ 1.8.0_31 mbed TLS (previously PolarSSL) 1.3.10 https://polarssl.org/ miTLS 0.1.3 http://www.mitls.org/ RSA BSAFE for C 4.0.4 http://www.emc.com/security/rsa-bsafe.htm RSA BSAFE for Java 6.1.1 http://www.emc.com/security/rsa-bsafe.htm Network Security Services (NSS) 3.17.4 https://developer.mozilla.org/en-US/docs/ Mozilla/Projects/NSS OpenSSL 1.0.1g https://www.openssl.org/ 1.0.1j 1.0.1l 1.0.2 nqsb-TLS 0.4.0 https://github.com/mirleft/ocaml-tls Table 1: Tested implementations unexpected behaviour, we perform a more in-depth anal- did not require client authentication, both are accept- ysis to determine the cause and severity. able paths. What is immediately clear is that there are An obvious first observation is that all the models more states than expected. Closer inspection reveals that of server-side implementations are very different. For there is a ‘shadow’ path, which is entered by sending example, note the huge difference between the mod- a HeartbeatRequest message during the handshake pro- els learned for RSA BSAFE for Java in Fig. 6 and for tocol. The handshake protocol then does proceed, but OpenSSL in Fig. 7. Because all the models are different, eventually results in a fatal alert (‘Internal error’) in re- they provide a unique fingerprint of each implementa- sponse to the Finished message (from state 8). From ev- tion, which could be used to remotely identify the imple- ery state in the handshake protocol it is possible to go to mentation that a particular server is using. a corresponding state in the ‘shadow’ path by sending the Most demo applications close the connection after HeartbeatRequest message. This behaviour is introduced their first response to application data. In the models by a security bug, which we will discuss below. Addi- there is then only one ApplicationData transition where tionally there is a redundant state 5, which is reached application data is exchanged instead of the expected cy- from states 3 and 9 when a ClientHello message is sent. cle consisting of an ApplicationData transition that al- From state 5 a fatal alert is given to all subsequent mes- lows server and client to continue exchanging application sages that are sent. One would expect to already receive data after a successful handshake. an error message in response to the ClientHello message In the subsections below we discuss the peculiarities itself. of models we learned, and the flaws they revealed. Cor- rect paths leading to an exchange of application data are Forgetting the buffer in a heartbeat As mentioned indicated by thick green transitions in the models. If there is any additional path leading to the exchange of above, HeartbeatRequest messages are not just ignored application data this is a security flaw and indicated by a in the handshake protocol but cause some side effect: dashed red transition. sending a HeartbeatRequest during the handshake proto- col will cause the implementation to return an alert mes- sage in response to the Finished message that terminates 4.1 GnuTLS the handshake. Further inspection of the code revealed the cause: the implementation uses a buffer to collect Fig. 2 shows the model that was learned for GnuTLS all handshake messages in order to compute a hash over 3.3.8. In this model there are two paths leading to a these messages when the handshake is completed, but successful exchange of application data: the regular one this buffer is reset upon receiving the heartbeat message. without client authentication and one where an empty The alert is then sent because the hashes computed by client certificate is sent during the handshake. As we server and client no longer match. 5 USENIX Association 24th USENIX Security Symposium 197 Figure 2: Learned state machine model for GnuTLS 3.3.8 Figure 3: Learned state machine model for GnuTLS 3.3.12. A comparison with the model for GnuTLS 3.3.8 in Fig. 2 shows that the superflous states (8, 9, 10, and 11) are now gone, confirming that the code has been improved. 6 198 24th USENIX Security Symposium USENIX Association #membership queries #equivalence queries Time (h:mm) Timeout #states GnuTLS 3.3.8 12 100ms 0:45 1370 5613 GnuTLS 3.3.12 7 100ms 0:09 456 1347 mbed TLS 1.3.10 8 100ms 0:39 520 2939 OpenSSL 1.0.1g + 16 100ms 0:31 1016 4171 OpenSSL 1.0.1j + 11 100ms 0:16 680 2348 OpenSSL 1.0.1l + 10 100ms 0:14 624 2249 OpenSSL 1.0.2 + 7 100ms 0:06 350 902 JSSE 1.8.0_25 9 200ms 0:41 584 2458 JSSE 1.8.0_31 9 200ms 0:39 584 2176 miTLS 0.1.3 6 1500ms 0:53 392 517 NSS 3.17.4 8 500ms 3:16 520 5329 RSA BSAFE for Java 6.1.1 6 500ms 0:18 392 517 RSA BSAFE for C 4.0.4 9 200ms 8:16 584 26353 nqsb-TLS 0.4.0 + 8 100ms 0:15 399 1835 + Without heartbeat extension Table 2: Results of the automated analysis of server implementations for the regular alphabet of inputs using our modified W-method with depth 2 Alphabet Algorithm Time (hh:mm) #states Membership queries Equivalence queries regular modified W-method 0:09 7 456 1347 full modified W-method 0:27 9 1573 4126 full original W-method 4:09 9 1573 68578 Table 3: Analysis of the GnuTLS 3.3.12 server using different alphabets and equivalence algorithms This bug can be exploited to effectively bypass the in- as at any time either one of the two parties is computing a tegrity check that relies on comparing the keyed hashes response, at which point it will not process any incoming of the messages in the handshake: when also resetting message. If an attacker would successfully succeed to this buffer on the client side (i.e. our test harness) at the exploit this issue no integrity would be provided on any same time we were able to successfully complete the message sent before, meaning a fallback attack would be handshake protocol, but then no integrity guarantee is possible, for example to an older TLS version or weaker provided on the previous handshake messages that were cipher suite. exchanged. By learning the state machine of a GnuTLS client we confirmed that the same problem exists when using 4.2 mbed TLS GnuTLS as a client. For mbed TLS, previously known as PolarSSL, we tested This problem was reported to the developers of version 1.3.10. We saw several paths leading to a suc- GnuTLS and is fixed in version 3.3.9. By learning mod- cessful exchange of data. Instead of sending a regular els of newer versions, we could confirm the issue is no ApplicationData message, it is possible to first send one longer present, as can be seen in Fig. 3. empty ApplicationData message after which it is still To exploit this problem both sides would need to reset possible to send the regular ApplicationData message. the buffer at the same time. This might be hard to achieve Sending two empty ApplicationData messages directly 7 USENIX Association 24th USENIX Security Symposium 199 after each other will close the connection. However, if in between these message an unexpected handshake mes- Client Server sage is sent, the connection will not be closed and only ClientHello a warning is returned. After this it is also still possible ServerHello; to send a regular ApplicationData message. While this is Certificate; strange behaviour, it does not seem to be exploitable. ServerHelloDone ClientKeyExchange; 4.3 Java Secure Socket Extension Finished For Java Secure Socket Extension we analysed Java ver- ChangeCipherSpec; sion 1.8.0_25. The model contains several paths leading {Finished} to a successful exchange of application data and contains more states than expected (see Fig. 4). This is the result ApplicationData of a security issue which we will discuss below. {ApplicationData} As long as no Finished message has been sent it is ap- parently possible to keep renegotiating. After sending a ClientKeyExchange, other ClientHello messages are ac- Figure 5: A protocol run triggering a bug in the JSSE, cepted as long as they are eventually followed by another causing the server to accept plaintext application data. ClientKeyExchange message. If no ClientKeyExchange message was sent since the last ChangeCipherSpec, a ChangeCipherSpec message will result in an error (state This issue was identified in parallel by Beurdouche et 7). Otherwise it either leads to an error state if sent di- al. [6], who also reported the same and a related issue for rectly after a ClientHello (state 8) or a successful change the client-side. By learning the client, we could confirm of keys after a ClientKeyExchange. that the issue was also present there. Moreover, after re- ceiving the ServerHello message, the client would accept the Finish message and start exchanging application data Accepting plaintext data More interesting is that the at any point during the handshake protocol. This makes model contains two paths leading to the exchange of ap- it possible to completely circumvent both server authen- plication data. One of these is a regular TLS protocol tication and the confidentiality and integrity of the data run, but in the second path the ChangeCipherSpec mes- being exchanged. sage from the client is omitted. Despite the server not receiving a ChangeCipherSpec message it still responds with a ChangeCipherSpec message to a plaintext Fin- 4.4 miTLS ished message by the client. As a result the server will MiTLS is a formally verified TLS implementation writ- send its data encrypted, but it expects data from the client ten in F# [8]. For miTLS 0.1.3, initially our test har- to be unencrypted. A similar problem occurs when trying ness had problems to successfully complete the hand- to negotiate new keys. By skipping the ChangeCipher- shake protocol and the responses seemed to be non- Spec message and just sending the Finished message the deterministic because sometimes a response was delayed server will start to use the new keys, whereas the client and appeared to be received in return to the next message. needs to continue to use its old keys. To solve this, the timeout had to be increased consider- This bug invalidates any assumption of integrity or ably when waiting for incoming messages to not miss confidentiality of data sent to the server, as it can be any message. This means that compared to the other im- tricked into accepting plaintext data. To exploit this issue plementations, miTLS was relatively slow in our setup. it is, for example, possible to include this behaviour in a Additionally, miTLS requires the Secure Renegotiation rogue library. As the attack is transparent to applications extension to be enabled in the ClientHello message. The using the connection, both the client and server applica- learned model looks very clean with only one path lead- tion would think they talk on a secure connection, where ing to an exchange of application data and does not con- in reality anyone on the line could read the client’s data tain more states than expected. and tamper with it. Fig. 5 shows a protocol run where this bug is triggered. The bug was report to Oracle and is 4.5 RSA BSAFE for C identified by CVE-2014-6593. A fix was released in their Critical Security Update in January 2015. By analysing The RSA BSAFE for C 4.0.4 library resulted in a model JSSE version 1.8.0_31 we are able to confirm the issue containing two paths leading to the exchange application was indeed fixed. data. The only difference between the paths is that an 8 200 24th USENIX Security Symposium USENIX Association Figure 4: Learned state machine model for JSSE 1.8.0_25 empty ApplicationData is sent in the second path. How- method and the analysis therefore takes much longer than ever, the alerts that are sent are not very consistent as they for the other implementations. differ depending on the state and message. For exam- ple, sending a ChangeCipherSpec message after an ini- tial ClientHello results in a fatal alert with reason ‘Ille- gal parameter’, whereas application data results in a fatal alert with ‘Unexpected message’ as reason. More cu- 4.6 RSA BSAFE for Java rious however is a fatal alert ‘Bad record MAC’ that is returned to certain messages after the server received the ChangeCipherSpec in a regular handshake. As this alert The model for RSA BSAFE for Java 6.1.1 library looks is only returned in response to certain messages, while very clean, as can be seen in Fig. 6. The model again other messages are answered with an ‘Unexpected mes- contains only one path leading to an exchange of appli- sage’ alert, the server is apparently able to successfully cation data and no more states than necessary. In gen- decrypt and check the MAC on messages. Still, an error eral all received alerts are ‘Unexpected message’. The is returned that it is not able to do this. This seems to be only exception is when a ClientHello is sent after a suc- a non-compliant usage of alert messages. cessful handshake, in which case a ‘Handshake failure’ is given. This makes sense as the ClientHello message is At the end of the protocol the implementation does not correctly formatted for secure renegotiation, which is not close the connection. This means we cannot take any required in this case. This model is the simplest that we advantage from a closed connection in our modified W- learned during our research. 9 USENIX Association 24th USENIX Security Symposium 201 Figure 6: Learned state machine model for RSA BSAFE for Java 6.1.1 4.7 Network Security Services Hello, whereas our test harness does this only on initial- isation of the connection. Therefore, the hash computed The model for NSS that was learned for version 3.17.4 by our test harness at the end of the handshake is not ac- looks pretty clean, although there is one more state than cepted and the Finished message in state 9 is responded one would expect. There is only one path leading to a to with an alert. Which messages are included in the hash successful exchange of application data. In general all differs per implementation: for JSSE all handshake mes- messages received in states where they are not expected sages since the beginning of the connection are included. are responded to with a fatal alert (‘Unexpected mes- sage’). Exceptions to this are the Finished and Heart- beat messages: these are ignored and the connection is closed without any alert. Other exceptions are non- handshake messages sent before the first ClientHello: Re-using keys In state 8 we see some unexpected be- then the server goes into a state where the connection haviour. After successfully completing a handshake, it is stays open but nothing happens anymore. Although the possible to send an additional ChangeCipherSpec mes- TLS specification does not explicitly specify what to sage after which all messages are responded to with a do in this case, one would expect the connection to be ‘Bad record MAC’ alert. This usually is an indication of closed, especially since it’s not possible to recover from wrong keys being used. Closer inspection revealed that this. Because the connection is not actually closed in this at this point OpenSSL changes the keys that the client case the analysis takes longer, as we have less advantage uses to encrypt and MAC messages to the server keys. of our modification of the W-method to decide equiva- This means that in both directions the same keys are used lence. from this point. We observed the following behaviour after the addi- 4.8 OpenSSL tional ChangeCipherSpec message. First, OpenSSL ex- pects a ClientHello message (instead of a Finished mes- Fig. 7 shows the model inferred for OpenSSL 1.01j. In sage as one would expect). This ClientHello is responded the first run of the analysis it turned out that Heartbeat- to with the ServerHello, ChangeCipherSpec and Fin- Request message sent during the handshake phase were ished messages. OpenSSL does change the server keys ‘saved up’ and only responded to after the handshake then, but does not use the new randoms from the Client- phase was finished. As this results in infinite models we Hello and ServerHello to compute new keys. Instead the had to remove the heartbeat messages from the input al- old keys are used and the cipher is thus basically reset phabet. This model obtained contains quite a few more (i.e. the original IVs are set and the MAC counter reset states than expected, but does only contain one path to to 0). After receiving the ClientHello message, the server successfully exchange application data. does expect the Finished message, which contains the The model shows that it is possible to start by sending keyed hash over the messages since the second Client- two ClientHello messages, but not more. After the sec- Hello and does make use of the new client and server ond ClientHello message there is no path to a successful randoms. After this, application data can be send over exchange of application data in the model. This is due the connection, where the same keys are used in both di- to the fact that OpenSSL resets the buffer containing the rections. The issue was reported to the OpenSSL team handshake messages every time when sending a Client- and was fixed in version 1.0.1k. 10 202 24th USENIX Security Symposium USENIX Association Figure 7: Learned state machine model for OpenSSL 1.0.1j Figure 8: Learned state machine model for OpenSSL 1.0.1g, an older version of OpenSSL which had a known security flaw [27]. 11 USENIX Association 24th USENIX Security Symposium 203 Early ChangeCipherSpec The state machine model would remove room for interpretation. of the older version OpenSSL 1.0.1g (Fig. 8) reveals a known vulnerability that was recently discovered [27], which makes it possible for an attacker to easily com- 5 Conclusion pute the session keys that are used in the versions up to 1.0.0l and 1.0.1g, as described below. We presented a thorough analysis of commonly used As soon as a ChangeCipherSpec message is received, TLS implementations using the systematic approach we the keys are computed. However, this also happened call protocol state fuzzing: we use state machine learn- when no ClientKeyExchange was sent yet, in which case ing, which relies only on black box testing, to infer a an empty master secret is used. This results in keys that state machine and then we perform a manual analysis of are computed based on only public data. In version 1.0.1 the state machines obtained. We demonstrated that this it is possible to completely hijack a session by sending is a powerful and fast technique to reveal security flaws: an early ChangeCipherSpec message to both the server in 3 out of 9 tested implementations we discovered new and client, as in this version the empty master secret is flaws. We applied the method on both server- and client- also used in the computation of the hash in the Finished side implementations. By using our modified version of message. In the model of OpenSSL version 1.0.1g in the W-method we are able to drastically reduce the num- Fig. 8 it is clear that if a ChangeCipherSpec message is ber of equivalence queries used, which in turn results in received too early, the Finished message is still accepted a much lower running time of the analysis. as a ChangeCipherSpec is returned (see path 0, 1, 6, 9, 12 Our approach is able to find mistakes in the logic in in the model). This is an indication of the bug and would the state machine of implementations. Deliberate back- be reason for closer inspection. The incoming messages doors, that are for example triggered by sending a par- after this path cannot be decrypted anymore however, be- ticular message 100 times, would not be detected. Also cause the corresponding keys are only computed by our mistakes in, for example, the parsing of messages or cer- test harness as soon as the ChangeCipherSpec message is tificates would not be detected. received, which means that these keys are actually based An overview of different approaches to prevent secu- on the ClientKeyExchange message. A simple modifi- rity bugs and more generally improve the security of soft- cation of the test harness to change the point at which ware is given in [38] (using the Heartbleed bug as a ba- the keys are computed will even provide a successful ex- sis). The method presented in this paper would not have ploitation of the bug. detected the Heartbleed bug, but we believe it makes a An interesting observation regarding the evolution of useful addition to the approaches discussed in [38]. It the OpenSSL code is that for the four different versions is related to some of the approaches listed there; in par- that we analysed (1.0.1g, 1.0.1j, 1.0.1l and 1.0.2) the ticular, state machine learning involves a form of neg- number of states reduces with every version. For ver- ative testing: the tests carried out during the state ma- sion 1.0.2 there is still one state more than required, but chine learning include many negative tests, namely those this is an error state from which all messages result in a where messages are sent in unexpected orders, which one closed connection. would expect to result in the closing of the connection (and which probably should result in closing of the con- nection, to be on the safe side). By sending messages in 4.9 nqsb-TLS an unexpected order we get a high coverage of the code, A recent TLS implementation, nqsb-TLS, is intended to which is different from for example full branch code cov- be both a specification and usable implementation writ- erage, as we trigger many different paths through the ten in OCaml [25]. For nsqb-TLS we analysed ver- code. sion 0.4.0. Our analysis revealed a bug in this imple- In parallel with our research Beurdouche et al. [6] in- mentation: alert messages are not encrypted even af- dependently performed closely related research. They ter a ChangeCipherSpec is received. This bug was re- also analyse protocol state machines of TLS implemen- ported to the nqsb-TLS developers and is fixed in a newer tations and successfully find numerous security flaws. version. What is more interesting is a design decision Both approaches have independently come up with the with regard to the state machine: after the client sends same fundamental idea, namely that protocol state ma- a ChangeCipherSpec, the server immediately responds chines are a great formalism to systematically analyse with a ChangeCipherSpec. This is different compared to implementations of security protocols. Both approaches all other implementations, that first wait for the client to require the construction of a framework to send arbi- also send a Finished message before sending a response. trary TLS messages, and both approaches reveal that This is a clear example where the TLS specifications are OpenSSL and JSSE have the most (over)complicated not completely unambiguous and adding a state machine state machines. 12 204 24th USENIX Security Symposium USENIX Association The approach of Beurdouche et al. is different though: provided to allow analysis of implementations. whereas we infer the state machines from the code with- The first manual analysis of the state machines we ob- out prior knowledge, they start with a manually con- tain is fairly straightforward: any superfluous strange be- structed reference protocol state machine, and subse- haviour is easy to spot visually. This step could even be quently use this as a basis to test TLS implementations. automated as well by providing a correct reference state Moreover, the testing they do here is not truly random, as machine. A state machine that we consider to be correct the ‘blind’ learning by LearnLib is, but uses a set of test would be the one that we learned for RSA BSAFE for traces that is automatically generated using some heuris- Java. tics. Deciding whether any superfluous behaviour is ex- The difference in the issues identified by Beurdouche ploitable is the hardest part of the manual analysis, but et al. and us can partly be explained by the difference for security protocols it makes sense to simply require in functionality that is supported by the test frameworks that there should not be any superfluous behaviour what- used. For example, our framework supports the Heart- soever. beat extension, whereas theirs supports Diffie-Hellman The difference behaviour between the various imple- certificates and export cipher suites. Another reason is mentations might be traced back to Postel’s Law: the fact that our approach has a higher coverage due to its ‘blind’ nature. ‘Be conservative in what you send, One advantage of our approach is that we don’t have to be liberal in what you accept.’ construct a correct reference model by hand beforehand. But in the end, we do have to decide which behaviour As has been noted many times before, e.g. in [35], this is unwanted. Having a visual model helps here, as it is is an unwanted and risky approach in security protocols: easy to see if there are states or transitions that seem re- if there is any suspicion about inputs they should be dis- dundant and don’t occur in other models. Note that both carded, connections should be closed, and no response approaches ultimately rely on a manual analysis to as- should be given that could possibly aid an attacker. To sess the security impact of any protocol behaviour that is quote [21]: ‘It’s time to deprecate Jon Postel’s dictum deemed to be deviant or superfluous. and to be conservative in what you accept’. When it comes to implementing TLS, the specifica- Of course, ideally state machines would be included in tions leave the developer quite some freedom as how the official specifications of protocols to begin with. This to implement the protocol, especially in handling errors would provide a more fundamental solution to remove – or exceptions. Indeed, many of the differences between or at least reduce – some of the implementation freedom. models we infer are variations in error messages. 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