SSH and TLS|Public-Key Encryption


We now look at two encryption mechanisms in popular use on the Internet. The first is the Secure Shell, SSH, used to allow login to remote systems, remote command execution, file transfer and even some forms of VPN tunneling. Public-key-encryption is optional; if it is used, the public keys are generally transported manually.

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The second example is the Transport Layer Security protocol, TLS, which is a successor of the Secure Sockets Layer, or SSL. Many people still refer to TLS by the SSL name even though TLS replaced SSL in 1999, though, to be fair, TLS is effectively an update of SSL, and TLS v1.0 could easily have been named SSL v4.0. TLS is used to encrypt web traffic for the HTTP Secure protocol, https; it is also used to encrypt traffic for several other applications and can be used, with appropriate programming, for any application.

If Alice wants to contact a server S using either SSH or TLS, at some point she will have to trust a public key claimed to be from S. A common approach with SSH is for Alice to accept the key on faith the first time it is presented, but then to save it for all future verifications. Under TLS, the key from S that Alice receives will have been signed by a certificate authority ( Certificate Authorities); Alice presumably trusts this certificate authority. (SSH does now support certificate authorities too, but their use with SSH seems not yet to be common.)

Both SSH and TLS eventually end up negotiating a shared-secret session key, which is then used for most of the actual data encryption.


The SSH protocol establishes an encrypted communications channel between two hosts, after establishing the identities of each endpoint. Its primary use is to enable secure remote-command execution with input and output via the secure channel; this includes the remote execution of an interactive shell, which is in effect a telnet-style terminal login with encryption. The companion program scp uses the SSH protocol to implement secure file transfer. Finally, ssh supports secure port forwarding from one machine (and port) to another; unrelated applications can then connect to one machine and find themselves securely talking to another. The current version of the SSH protocol is 2.0, and is defined in RFC 4251. The authentication and transport sub-protocols are defined in RFC 4252 and RFC 4253 respectively.

One of the first steps in an SSH connection is for the endpoints to negotiate which secret-key cipher (and mode) to use. Support of the following ciphers is “recommended” and there is a much longer list of “optional” ciphers (which include RC4 and Blowfish); the table below includes those added by RFC 4344:

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SSH supports a special name format for including new ciphers for local use.

The SSH protocol also allows the endpoints to negotiate a public-key-encryption mechanism, a secure-hash function, and even a key-exchange algorithm although only minor variants of Diffie-Hellman-Merkle key exchange are implemented.

If Alice wishes to connect to a server S, the server clearly wants to verify Alice’s identity, either through a password or some other means. But it is just as important for Alice to verify the identity of S, to prevent man-in-the-middle attacks and the possibility that an attacker masquerading as S is simply collecting Alice’s password for later use.

The SSH protocol is not designed to minimize the number of round-trip packet exchanges, making SSH connection setup quite a bit slower than TLS connection setup ( TLS Connection Setup). A connection may involve quite a few round trips to get started: the TCP three-way handshake, the protocol version exchange, the “Key Exchange Init” exchange, the actual Diffie-Hellman-Merkle key exchange, the “NewKeys” exchange, and the “Service Request” exchange (all but the first are documented in RFC 4253). Still, multi-second delays can usually be reduced through performance tuning; enabling diagnostic output (eg with the -v option) is often helpful. A common delay culprit is a server-side DNS lookup of the client, fixable with a server-side configuration setting of UseDNS no or the equivalent.

In the following subsections we focus on the “common” SSH configuration and ignore some advanced options.

Server Authentication

To this end, one of the first steps in SSH connection negotiation is for the server to send the public half of its host key to Alice. Alice verifies this key, which is typically in her known_hosts file. Alice also asks

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S to sign something with its host key. If Alice can then decrypt this with the public host key of S, she is confident that she is talking to the real S.

If this is Alice’s first attempt to connect to S, however, she should get a message like the one below:

The authenticity of host ‘S (’ can’t be established.

RSA key fingerprint is da:2e:e3:94:84:6b:bf:6d:2f:e4:c3:76:68:72:a5:a0.

Are you sure you want to continue connecting (yes/no)?

If Alice is cautious, she may contact the administrator of S and verify the key fingerprint. More likely, she will simply choose “yes”, in which case the host key of S will be added to her own known_hosts file; this latter strategy is sometimes referred to as trust on first use. If she later re-connects to S after the host key of S has been changed, she will get a rather more dire message, and, under the default configuration, she will not be allowed to continue until she manually removes the old, now-incorrect, host key for S from her known_hosts file.

See also the ARP-spoofing scenario at 10.2.2 ARP Security, and the comment there about how this applies to SSH.

SSH v2.0 now also supports a certificate-authority mechanism for verifying server host keys, replacing the known_hosts file.

Key Exchange

The next step is for Alice’s computer and S to negotiate a session key. After the cipher and key-exchange algorithms are negotiated, Alice’s computer and S use the chosen key-exchange algorithm to agree on a session key for the chosen cipher.

The two directions of the connection generally get different session keys. They can even use different encryption algorithms.

At this point the channel is encrypted.

Client Authentication 

The server now asks Alice to authenticate herself. If password authentication is used, Alice types in her password and it and her username are sent to the server, over the now-encrypted connection. This does expose the server to brute-force password-guessing attacks, and is not infrequently disabled.

Alice may also have set up RSA authentication (other types of public-key authentication are also possible). For this, Alice must create a public/private key pair (often in files and id_rsa), and the public key must have been previously installed on S. On Unix-based systems it is often installed on S in the authorized_keys file in the .ssh subdirectory of Alice’s home directory. If Alice now sends a message to S signed by her private key, S is in a position to verify the signature. If this succeeds, Alice can log in without supplying a password to S.

It is common though not universal practice for Alice’s private-key file (on her own computer) to be protected by a password. If this is the case, Alice will need to supply that password, but it is used only on Alice’s end of the connection. (The default password hash is MD5, which may not be a good choice; see 28.6.2 Password Hashes.)

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Public-key authentication is tried before password authentication. If Alice has created a public key, it is likely to be tried even if she has not copied it to S.

If S is set up to require public-key authentication, Alice may not have any direct way to install her public key on S herself, and may need the cooperation of the administrators of S. It is possible that the latter will send Alice a public/private keypair chosen by them specifically for S, rather than allowing Alice to choose her own keypair. The standard SSH user configuration does support different private keys for different servers.

In several command-line implementations of ssh, the various stages of authentication can be observed from the client side by using the ssh or slogin command with the -v option.

The Session

Once an SSH connection has started, a new session key is periodically negotiated. RFC 4253 recommends this after one hour or after 1 GB of data is exchanged.

Data is then sent in packets generally with size a multiple of the block size of the cipher. If not enough data is available, eg because only a single keystroke (byte) is being sent, the packet is padded with random data as needed. Every packet is required to be padded with at least four bytes of random data to thwart attacks based on known plaintext/ciphertext pairs. Included in the encrypted part of the packet is a byte indicating the length of the padding.


Transport Layer Security, or TLS, is an IETF extension of the Secure Socket Layer (SSL) protocol originally developed by Netscape Communications. SSL went through published versions 2.0 and 3.0; the latter was introduced in 1996 and was officially deprecated by RFC 7568 in 2015 (see the POODLE sidebar above at 29.5 SSH and TLS). TLS 1.0 was introduced in 1999, as the IETF took ownership of the specification from Netscape. The current version of TLS is 1.2, specified in RFC 5246 (plus updates listed there). A draft of TLS 1.3 is in progress (2018); see draft-ietf-tls-tls13.

The original and still primary role for TLS is encrypting web connections and verifying for the client the authenticity of the server. TLS can, however, be embedded in any network application.

Unlike SSH, client authentication, while possible, is not common; web servers often have no pre-existing relationship with the client. Also unlike SSH, the public-key mechanisms are all based on established certificate authorities, or CAs, whereas the most common way for an SSH server’s host key to end up on a client is for it to have been accepted by the user during the first connection. Browsers (and other TLS applications as necessary) have embedded lists of certificate authorities, known as trust stores, trusted by the browser vendor. SSH requires no such centralized trust.

If Bob wishes to use TLS on his web server SBob, he must first arrange for a certificate authority, say CA1, to sign his certificate. A certificate contains the full DNS name of SBob, say, a public key KS used by SBob, and also an expiration date. TLS uses the ITU-T X.509 certificate format.

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Now imagine that Alice connect to Bob’s server SBob using TLS. Early in the process SBob will send her its signed certificate, claiming public key KS. Alice’s browser will note that the certificate is signed by CA1, and will look up CA1 on its list of trusted certificate authorities. If found, the next step is for the browser to use CA1’s public key, also on the list, to verify the signature on the certificate SBob sent.

If everything checks out, Alice’s browser now knows that CA1 certifies has public key KS. As SBob has presented this key KS, and is able to verify that it possesses the matching private key, this is proof that SBob is legitimately the server with domain name

Assuming, of course, that CA1 is correct. As with SSH,

once Alice has accepted the public key KS of SBob, a secret-key cipher is negotiated and the remainder of the exchange is encrypted.

Certificate Authorities

A certificate authority, or CA, is just an entity in the business of signing certificates: messages that include a name SBob and a verified public key KS. The purpose of the certificate authority is to prevent man-in-themiddle attacks (29.3 Trust and the Man in the Middle); Alice wants to be sure she is not really connected to SBad instead of to SBob.

What the CA is actually signing is that the particular public key KS belongs to domain, just like Charlie might sign Bob’s public key on an individual basis. The difference here is that the CA is likely to be a large organization, and the CA’s public key is likely to be embedded in the network application software somewhere. (Real CA’s usually have a two-layer key signing arrangement, in which Bob’s public key would be signed by one of the CA’s subsidiary keys, but the effect is the same.)

A certificate specific for will not work for, even though both DNS names might direct to the same server SBob. If SBob presents a certificate for to a browser that has arrived at SBob via the url, the user should see a warning. It is straightforward, however, for the server either to support two certificates, or (if the certificate authority supports this) a single certificate for with a Subject Alternative Name also entered on the certificate for alone. It is also possible to obtain “wildcarded” certificates for * (though this does not match; it also does not match names with two additional levels such as RFC 6125, §7.2, recommends against wildcard certificates, though they remain widely used.

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Popular browsers all use preset lists of CAs provided by (and presumably trusted by) either the browser vendor or the host operating-system vendor. If Alice is also willing to trust these CAs, she can feel comfortable

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using the key she receives to send private messages to Bob. That “if”, however, is sometimes controversial. Just because Alice trusts her browser and operating system (eg not to contain malware), that does not automatically imply that Alice should trust these vendors’ judgment when it comes to CAs.

On the face of it, Bob’s certificate authority CA1 is just signing that domain name has public key KS. This isn’t quite the same, however, as attesting that the certificate-signing request actually came from Bob. All depends on how thorough CA1 is in checking the identity of its customer, and since those customers typically choose the least expensive CA, there is sometimes an incentive to cut corners.

A certificate authority’s failure to verify the identity of the party making the certificate-signing request can enable a man-in-the-middle attack (29.3 Trust and the Man in the Middle). Suppose, for example, that Mallory is able to obtain a signed certificate from another certificate authority CA2 linking to key Kbad controlled by Mallory. If Mallory can now position his server SBad in the middle between Alice and SBob,

Alice SBad SBob

then Alice can be tricked into connecting to SBad instead of SBob. Alice will request a certificate, but from SBad instead of SBob, and get Mallory’s from CA2 instead of Bob’s actual certificate from CA1. Mallory opens a second connection from SBad to SBob (this is easy, as Bob makes no attempt to verify Alice’s identity), and forwards information from one connection to the other. As far as TLS is concerned everything checks out. From Alice’s perspective, Mallory’s false certificate vouches for the key Kbad of SBad, CA2 has signed this certificate, and CA2 is trusted by Alice’s browser. There is no “search” at any point through the other CAs to see if any of them have any contrary information about SBob. In fact, there is not necessarily even contact with CA2, though see Certificate revocation below.

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If the certificate authority CA1 were also the domain registrar with whom Bob registered the DNS name, it would be especially well-positioned to verify that Bob is really the owner of the domain. But this is not generally the case. Quite often, the CA’s primary verification method is to send an email to, say, (or perhaps to the DNS administrative contact listed in the WHOIS database for domain If someone responds to this email, it is assumed they must be legitimately part of the domain. Another authentication strategy is for CA1 to send a special file to be placed at, say,; again, the idea is that only the domain owner would be able to do this. Note, however, that in the man-in-the-middle attack above, it does not matter what CA1 does; what matters is the verification policy used by CA2.

If Alice is very careful, she may click on the “lock” icon in her browser and see that the certificate authority signing her connection to SBad is CA2 rather than CA1. But if Alice has a secure way of finding Bob’s “real” certificate authority, she might as well use it to find Bob’s key KS. As she often does not, this is of limited utility in practice.

The second certificate authority CA2 might be a legitimate certificate authority that has been tricked, coerced or bribed into signing Mallory’s certificate. Alternatively, it might be Mallory’s own creation, inserted by

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Mallory into Alice’s browser through some other vulnerability

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Mallory’s machinations here do require both the man-in-the-middle attack and the bad certificate. If Alice is able to establish a direct connection with SBob, then the latter will send its true key KS signed by CA1.

As another attack, Mallory might obtain a certificate for and hope Alice doesn’t notice the spelling difference between B0B and BOB. When this is done, Mallory often also sends Alice a disguised link to in the hope she will click on it. Unicode domain names make this even easier, as Unicode provides many character pairs that are different but which look identical.

Certificates tend to come with a number of constraints. First among them is that most certificates issued by CAs to end users (either corporations or individuals) are marked as “not a certificate authority”, meaning that the certificate cannot be used to sign additional certificates. That is, the certificate trust relationship is not transitive. If Bob gets a certificate for, he cannot use it to sign a certificate for Alice or, for that matter, Google. The second major restriction is on names: a certificate is almost always tied to a specific DNS name. As described above, multiple names can be accepted via the “subject alternative name” attribute, or by using a “wildcard” name like * Finally, certificates often contain restrictions on how they can be used, in the “extended key usage” attribute. A typical entry is “TLS web server authentication”, expressed as an OID (26.3 SNMP Naming and OIDs). These restrictions are enforced by the TLS library; it is possible outside that library to use the keys inside certificates for whatever you want.

Extended-Validation certificates were introduced in 2007 as a way of providing greater assurances that the certificate issued to was in fact generated by a request from Bob himself; Mallory should in theory have a much harder time obtaining an EV certificate for from CA2. Desktop browsers that have secured a TLS connection using an EV certificate typically add the name of the domain owner, highlighted in green and/or with a green padlock icon, to the address bar. Financial institutions often use EV certificates. So does

Of course, if Alice does not know Bob is using an EV certificate, she can still be tricked by Mallory as above. Alternatively, perhaps Alice is using a mobile device; most mobile browsers give little if any visual indication of EV certificates. Because the majority of browsing today (2018) is done on mobile devices, this severely limits the usefulness of these certificates.

Certificate pinning

Another strategy intended as a more direct prevention of man-in-the-middle attacks is certificate pinning. Conceptually, Alice (or her browser) makes a note the first time she connects to SBob that the certificate authority is CA1 or that Bob’s public key is KS, or both. Future connections to SBob must match at least one of these credentials. This form of pinning depends on Alice’s having reached the real SBob on the first connection, sometimes called “trust on first use”. A similar trust-on-first-use strategy is often (though not always) used with SSH, Server Authentication.

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In the pinning protocol described in RFC 7469, it is SBob that initiates the pinning by including a “pin directive” in its initial HTTP connection. This requests Alice’s browser to pin the desired certificates, though the pinned correspondence between a site and its keys is always maintained at the browser side. The pin directive also specifies which keys (KS or CA1 or both) are pinned, and for how long.

If SBob sends a pin directive for KS, then Alice’s browser remembers KS, and any new certificate at SBob will be a mismatch. If SBob pins CA1, then SBob can have CA1 issue a new key and it will still pass pinvalidation. If CA1 is later compromised, though, SBob is not protected against any new rogue certificates it issues. (This is generally a minor concern, as CA1 compromise will always mean that new visitors to SBob will be vulnerable to man-in-the-middle attacks.)

Bob can also obtain, along with KS, a backup certificate KS-backup, and have SBob’s pin directives pin both of these. Then, if KS is compromised, KS-backup is ready to go. Hopefully, key compromise is a rare event, so there is a very small chance that both KS and KS-backup are compromised within the pin’s lifetime. Key compromise pretty much follows from any server compromise, however, and if intruders break in, KS-backup cannot be put into production until Bob is sure the site is secure. That may take some time.

If KS and KS-backup are both compromised, and CA1 wasn’t pinned, SBob may simply become inaccessible to returning visitors. Automatic unpinning of revoked certificates would help, but certificate revocation (see following section) has its own difficulties. The user may have an option to disable pin validation for this particular site, but it’s not supposed to be as simple as clicking “ok”, and in any event if the site is your bank you may be loath to do this.

Certificate revocation

Suppose the key r underlying the certificate for has been compromised, and Mallory has the private key. Then, if Mallory can trick Alice into connecting to SBad instead of SBob, the original CA1 will validate the connection. SBad can prove to Alice that it has the secret key corresponding to KS, and all the certificate does is to attest that has key KS. Mallory’s deception works even if Bob noticed the compromise and updated SBob’s key to K2; the point is that Mallory still has the original key, and CA1’s certificate attesting to that original key is still valid.

There is a mechanism by which a certificate can be revoked. Revocation information, however, must be kept at some central directory; a server can continue to serve up a revoked certificate and unless the clients actively check, they will be none the wiser. This is the reason certificates have expiration dates.

The original revocation mechanism was the global certificate revocation list. A newer alternative is the Online Certificate Status Protocol, OCSP, described in RFC 6960. If Alice receives a certificate signed by CA1, she can send the serial number of the certificate to a designated “OCSP responder” run by or on behalf of CA1. If the certificate is still valid, the responder site will return a signed confirmation of that.

Of course, an eavesdropper watching Alice’s traffic arriving at the OCSP responder – and the OCSP responder itself – now knows that Alice is visiting An eavedropper closer to Alice, however, knows that anyway.

More seriously, someone running a man-in-the-middle attack close to Alice can probably intercept and block Alice’s OCSP request. If Alice receives no response, what should she do? Maybe there is a problem, but maybe the responder site is simply down or the Internet is simply slow. Most browsers that actually do revocation checks assume the latter – known as soft fail – making revocation checks of dubious value. The alternative of refusing to allow access to the original site – hard fail – leads to many false positives.

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As of 2016, there is no generally accepted solution. One minimal approach, in the event of OCSP soft fail, is to use hard fail with EV certificates, or at least to downgrade EV certificates to ordinary ones upon OCSP non-response. Another approach is OCSP stapling, in which the server site periodically (perhaps daily) requests a signed and dated update from its CA’s OCSP server indicating that the site’s certificate is still valid. This OCSP report is then “stapled” to the ServerHello message (below), if requested by the client. This allows the client to verify that the certificate was still valid quite recently.

Of course, if the client is the victim of a man-in-the-middle attack then the (fake) server will not staple an OCSP validity report, and the client must fall back to the regular OCSP lookup process. But this case can be expected to be infrequent, making a hard fail after OCSP non-response a reasonable option.

Google’s Chrome browser implements CRLSets in lieu of checking OCSP servers, described here. This is a list of revoked certificates downloaded regularly to the browser. Unfortunately, the full list is much too large, so CRLSets are limited to emergency revocations and certificate revocations due to key compromise; even then the list is not complete.

TLS Connection Setup

The typical TLS client-side user is interested in viewing a web page as quickly as possible, placing a premium on rapid negotiation of the TLS connection. If sites were to load noticeably more slowly when encryption was used, encryption might not be used routinely. As a result, the connection-setup process, known as the TLS handshake protocol, is designed to complete in two RTTs. The goals of the handshake protocol are to agree on the encryption and authentication mechanisms to be used, to provide authentication as necessary, and to negotiate the encryption keys. Here is an outline of a typical exchange, with some options omitted:

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The client initiates the connection by sending its ClientHello message, which contains its preferred TLS/SSL-protocol version number, a session identifier, some random bytes (a cryptographic nonce), and a list of cipher suites: tuples consisting of a key-exchange algorithm, a shared-key encryption algorithm and a secure-hash algorithm. The list of cipher suites is ranked by the client according to preference.

Many large servers host multiple websites through “virtual hosting”. Each website will have its own certificate, and the server needs to be able to know which one to send in its reply. In settings like this, the client includes the server’s DNS name in a ServerName extension; the Server Name Indication (SNI) is defined in RFC 6066.

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The server then responds with its ServerHello. This contains the final protocol-version number, generally the maximum of the version proposed by the client and the highest version supported by the server. The ServerHello also contains the server’s choice of a cipher suite from the client’s list, and some more random bytes (the server’s nonce). The server also sends its certificate, if requested (which it almost always is). Exactly which certificate is sent may depend on the server name sent by the client in the optional ServerName extension. The certificate is considered to be a separate “message” at the TLS protocol level, not part

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of the ServerHello, but everything is sent together.

Having received the server’s certificate, the client application’s next step is to validate this certificate, by checking its signature against the client’s list of trusted certificate authorities, the client’s “trust store”. This step does not involve any network communication. While most operating systems maintain a list of generally trusted certificates, the client application can trust all, some or none of this list; the client can also load application-specific certificates. If the certificate does not check out, the client is free to continue the connection, perhaps pausing to add the server certificate to its trust store (hopefully after user interaction confirming this).

The client then responds with its key-exchange response, and its own certificate if applicable (which it seldom is). It immediately follows that with its Finished message, the first message that is encrypted with the just-negotiated cipher suite. The server then replies with its encrypted Finished message, and encrypted application-layer communication can then begin.

Most browsers allow the user to click on the padlock icon for the TLS-secured connection to find out what cipher suite was actually agreed upon.

A client starting a brand-new connection leaves the Session ID field empty in the ClientHello message. However, if a client wishes to resume a previous session, it includes here the Session ID from that previous session. The server must, of course, also recognize that session.

Once upon a time, some broken TLS servers failed to respond properly if a client proposed a version number greater than what the server supported; the server would close the connection instead of returning a lower version number. As a result, many clients were programmed to try again with the next-lower version in the event of any connection-setup failure. This meant that an attacker could force a client to propose an obsolete version (eg SSL 3.0) simply by interrupting earlier connection-setup attempts, perhaps with a TCP RST packet. A consequence was the POODLE vulnerability mentioned in the sidebar in 29.5 SSH and TLS.

An application’s choice of TLS (versus unencrypted communication) is often signaled by the use of a special port number; for example, the standard http port is 80 while the standard https port is 443. Alternatively, there may be some initial negotiation; for example, in the SMTP email protocol (RFC 5321) it is common for the client to connect to the server on the standard port 25 without encryption, but then to negotiate the use of TLS using the STARTTLS extension to SMTP (RFC 3207). This adds multiple extra RTTs, but for non-interactive protocols this is usually of only minor concern.

Occasionally applications do sometimes get confused about whether TLS is in use. Web browsers connecting to port 80 (HTTP) or port 443 (HTTPS) can usually figure things out correctly, even if the “http://” or “https://” part of the URL is missing. However, browsers connecting to nonstandard webservers running on nonstandard ports may receive a very cryptic binary response if TLS was required but the “https://” URL prefix was missing.

Domain Fronting

Government censorship of websites and other Internet services – such as encrypted messaging – is rampant in many parts of the world. We saw above that the ServerName extension is included in many ClientHello messages in order to request the correct certificate. This is sent in the clear, and can be used by censors to block selected traffic.

One censorship workaround, known as domain fronting, is to ask for one hostname, say, in the ClientHello ServerName extension, but then request a second, different hostname,

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within the application-layer protocol, after encryption is in place. For example, an HTTP GET request (17.6.2 netcat again) almost always contains a HOST: header; this is where would be specified. Of course, this only works if and are hosted at the same IP address, but large HTTP servers often do handle multiple websites on a single server using a mechanism known as virtual hosting (10.1.2 nslookup and dig). To block, a censor must now also block The hope is that censors will be unwilling to do that.

Domain fronting is even easier when CDNs (1.12.2 Content-Distribution Networks) are involved; virtual hosting is not necessary. A CDN accepts connections to the websites of a large number of customers (often a much larger number than a single server can support using virtual hosting) a each edge server of their network. The edge servers have certificates for each of the customers, and so appear legitimate to users. After the connection, the edge server might then answer some requests with cached data, but will forward other requests on to the actual web server. To enable domain fronting in this environment, might only need to become a customer of the same CDN as

CDNs and large hosting providers sometimes do not think highly of domain fronting, particularly as some censors have shown willingness to block critical domains like just to block an unwanted application. Google and Amazon Web Services disabled domain fronting for their CDN customers in 2018.

Before the initial connection to the server, the client will usually make a DNS request to obtain the server’s IP address. These too can be blocked, or monitored, although there are workarounds.

TLS key exchange

Perhaps the most important part of the TLS handshake is to negotiate the encyption keys. The keys themselves are all derived from the TLS master secret. The key derivations are done in deterministic ways (eg using secure hashes, or a stream-cipher algorithm), and can be done by either side once the master secret is known.

In all cases, the client should know the master secret after receiving the ServerHello and related packets. At this point – after one RTT – the client can begin sending encrypted messages to the server. At this point, the client can even begin sending encrypted data, although it is possible that the client is not yet fully authenticated to the server.

Negotiation of the master secret depends on the cipher suite chosen. In the RSA key-exchange method, the client chooses a random pre-master secret. This is sent to the server in the third leg of the exchange, after the client has received the server’s certificate in the ServerHello stage. The pre-master secret is encrypted with the public RSA key from the server’s certificate, thus conveying it securely to the server. Both sides calculate the master secret from the pre-master secret and both sides’ cryptographic nonces, using a secure hash or the pseudorandom keystream of a stream cipher (28.7.4 Stream Ciphers). The inclusion of the server nonce means tha the client does not unilaterally specify the key.

If Diffie-Hellman-Merkle key-exchange is used (28.8 Diffie-Hellman-Merkle Exchange), then the server proposes the prime p and primitive root g during the ServerHello phase, as part of its selection of one of the cipher suites listed in the ClientHello. The server (Bob) also includes gb mod p. The client chooses its exponent a, and sends ga mod p to the server. The pre-master secret is gaˆb mod p, which the client, as in the RSA case, knows at the end of the first RTT. The client’s ga mod p cannot be encrypted using the master key, but everything else sent by the client at that point can be. There is also a version of Diffie-Hellman-Merkle exchange that uses elliptic-curve cryptography.

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TLS version 1.3

Version 1.3 of TLS, RFC 8446, makes numerous improvements to the protocol. Among other things, it prunes the list of acceptable cipher suites of all algorithms no longer considered secure, and also of all keyexchange algorithms that do not support forward secrecy, 29.2 Forward Secrecy. The version-negotiation process has also been improved, to prevent version-downgrade attacks, and the 0-RTT mode (below) enables faster connection startup.

Of the two key-exchange methods in the previous section, Diffie-Hellman-Merkle does support forward secrecy while RSA may not, depending on the latter’s choice of the pre-master secret. The proposal to drop RSA key exchange has caused concern in some enterprise settings, particularly the banking industry, where widespread use is often made of middleboxes that are able to decrypt and monitor internal TLS connections (that is, connections where both endpoints are under control of the same enterprise). This monitoring is done for a variety of reasons, including regulatory compliance and malware detection.

To enable such monitoring using TLS v1.2, all TLS connections between the enterprise and the outside Internet terminate at the site’s border, at a special TLS-proxy appliance. From there, connections to interior systems are re-encrypted, but using a static RSA pre-master secret that is shared with the monitoring equipment. The same applies to any TLS connections that are entirely internal to the enterprise.

To enable this kind of monitoring under TLS v1.3, one approach is to have the monitoring occur at the TLS-proxy appliance, though this means that device will have to be considerably bigger, more complicated and more expensive; additionally, it will not help with entirely internal TLS connections. Another approach is to have the TLS proxy, and at least one endpoint of internal TLS connections, turn over all session keys to the monitors. This raises other security concerns. For two proposed amendments to the original TLS v1.3 proposal, see the Internet Drafts draft-rhrd-tls-tls13-visibility and draft-green-tls-static-dh-in-tls13.

TLS v1.3 requires that all messages after the ServerHello be encrypted (except possibly one last clientto-server key-exchange message that should be intrinsically secure). The client knows the master secret at this point, and the server will know it as soon as it hears from the client.

TLS v1.3 0-RTT mode

Finally, TLS v1.3 adds a so-called O-RTT mode, in which the client can send encrypted messages from the very beginning, assuming an appropriate master secret has been negotiated during a previous connection. For HTTPS interactions, this is a common case. The 0-RTT mode, in particular, allows the client to send secure data along with its ClientHello. The server can respond with 0-RTT-protected data along with its ServerHello, if it chooses to. After that, everything should be protected with the new session’s master secret; this is known as 1-RTT protection.

To send 0-RTT data, the client must resume the earlier session by including the previous Session ID in its ClientHello message. The server can refuse 0-RTT data, and demand 1-RTT protection, but usually will not.

The advantage of 0-RTT data is that, if the server accepts it, an answer can return to the client in a single RTT. This is particularly significant in QUIC (16.1.1 QUIC and 18.15.4 QUIC Revisited), in which the TLS handshake is carried out in parallel with the QUIC connection handshake (a replacement for the TCP three-way handshake). This means that the response comes just one RTT after the very first message from the client. QUIC requires TLS v1.3 (or later).

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0-RTT protection is not quite as strong ast 1-RTT protection. For one, forward secrecy does not apply. More seriously, a participant receiving 0-RTT data is vulnerable to replay attacks. An attacker cannot modify a previously intercepted 0-RTT message (without breaking the cipher), but can resend it.

At a minimum, the recipient of a 0-RTT request should accept it only if it is idempotent: yielding the same results and side-effects whether executed once or twice (16.5.2 Sun RPC). This is generally automatic for simple HTTP GET requests. Idempotency is not itself a security guarantee – after all, the request “delete all files” is idempotent. The point, however, is that the particular request must, if part of a replay attack, have been executed before, and if it was safe once, it should be safe twice. A TLS server can also suppport other anti-replay mechanisms, such as a database of past requests.

Another consequence of mixing 0-RTT and 1-RTT data is that the recipient needs to be able to tell which requests received 0-RTT protection and which received full 1-RTT protection.

A TLS Programming Example

In this section we introduce a simple pair of programs, tlsserver.c and tlsclient.c, that communicate via TLS. They somewhat resemble the simplex-talk program in 17.6 TCP simplex-talk, except that neither reads from the command line. Instead, the client sends a message to the server (which may ignore it), and then the server sends a message back. This structure will allow us later to point the client at a “real” TLS-using (that is, HTTPS-using) web server, instead of tlsserver, have the client send an HTTP GET request (17.6.2 netcat again), and obtain the web server’s response.

Both programs are written in C, in order to use the OpenSSL library, a descendant of the original Netscape SSL implementation. The openssl package also includes some important command-line utilities for certificate creation. While OpenSSL library documentation remains notoriously spare, parts of the library have now been audited, and OpenSSL remains the most widely used implementation of TLS.

The code shown here is intended simply as a demonstration; it should not be considered a model of how to implement secure connections.

Making a certificate

The first step is to create the appropriate application certificate, and create our own certificate authority to sign it. For each step we will use the openssl command, also used below at 29.8 RSA Key Examples with additional explanation. We start with the certificate authority. The first step is to create the key our certificate authority will use; this is what the genrsa option below achieves. We choose a key length of 4096 bits.

openssl genrsa -out CAkey.pem 4096


The resultant file says it is a “private key”, but it is in fact a public/private key pair.

The next step is to create a self-signed certificate. The req option asks for a signing “request”, the -new option indicates this is a new request, and the -x509 option tells openssl to forget making a “request” and instead make a self-signed certificate. X.509 is actually a format standard.

openssl req -new -x509 -key CAkey.pem -out CAcert.pem -days 10000


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The certificate here is set to expire in 10,000 days; real published certificates are not supposed to be valid for more than about three years. The signing process (self- or not) triggers a series of questions that will form the “Distinguished Name” of the certificate:

Country Name (2 letter code): US
State or Province Name (full name): Illinois
Locality Name (eg, city):Shabbona
Organization Name (eg, company): An Introduction to Computer Networks Organizational Unit Name (eg, section): Security
Common Name (e.g. server FQDN or YOUR name): Peter L Dordal Email Address []:


If we were requesting a certificate for a web server, we would use the hostname as Common Name, but we are not. The output here, CAcert.pem, represents the concatenation of the above information with the public key from CAkey.pem, and then signed by the private key from CAkey.pem. The private key itself, however, is not present in CAcert.pem; this is the file that the TLS client will receive as its certificate authority. We can read the information from the file as follows:

openssl x509 -in CAcert.pem -text


This produces the following, somewhat edited for space. We can see that the certificate is self-signed because the Issuer: is the same as the Subject:

Version: 3 (0x2)
Serial Number: 14576542390929842281 (0xca4a481320cbe069)
Signature Algorithm: sha256WithRSAEncryption
Issuer: C=US, ST=Illinois, L=Shabbona, O=An Introduction to Computer ãÑNetworks, OU=Security, CN=Peter L Dordal
Not Before: Jan 10 20:17:55 2017 GMT
Not After : May 28 20:17:55 2044 GMT
Subject: C=US, ST=Illinois, L=Shabbona, O=An Introduction to Computer ãÑNetworks, OU=Security, CN=Peter L Dordal
Subject Public Key Info:
Public Key Algorithm: rsaEncryption
Public-Key: (4096 bit)
00:b0:70:06:7e:38:1d:29:35:a7:ca:40:bf:fd:6e: e5:26:7b:ee:0d:e7:d7:c2:61:8e:42:5f:b9:85:8c:
Exponent: 65537 (0x10001)
X509v3 extensions:
X509v3 Subject Key Identifier: D1:74:15:F5:31:CB:DD:FA:D6:AE:81:7A:40:AA:64:7A:55:96:2E:08
X509v3 Authority Key Identifier: ãÑkeyid:D1:74:15:F5:31:CB:DD:FA:D6:AE:81:7A:40:AA:64:7A:55:96:2E:08 X509v3 Basic Constraints: CA:TRUE


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Signature Algorithm: sha256WithRSAEncryption 2a:d2:35:43:c2:5d:1c:5d:e2:88:ed:4e:aa:d2:b5:d6:e9:26: 60:f0:37:ea:29:56:14:62:58:01:78:b0:6f:ee:ab:40:17:36:



We now create the application key and then the application signature request. This request we will then sign with the above certificate CAfile.pem to generate the application certificate.

openssl genrsa -out appkey.pem 2048
openssl req -new -key appkey.pem -out appreq.pem -days 10000

The certificate request here again requires entering a Distinguished Name. This time we enter as follows; the Organization Name must match the CAcert.pem above:

Country Name (2 letter code): US
State or Province Name (full name): Illinois
Locality Name (eg, city): no fixed abode
Organization Name (eg, company): An Introduction to Computer Networks Organizational Unit Name (eg, section): TLS server
Common Name (e.g. server FQDN or YOUR name): Odradek
 Email Address []:


We are also asked for a “challenge password”, but we leave this blank.

Now we come to the final step: the actual signing. Unlike any of the previous openssl commands, this requires root privileges. It also makes use of the global openssl configuration file (/etc/ssl/openssl.cnf), which, among other things, references a file “serial” to assign the certificate a serial number. (We did not have to do any of this to create the self-signed certificate.)

openssl ca -in appreq.pem -out appcert.pem -days 10000 -cert CAcert.pem - ãÑkeyfile CAkey.pem


We now have our application certificate in appcert.pem, which we deliver to the application section, below. If we read it with openssl x509 -in appcert.pem -text, we find that the Issuer is that of CAfile.pem, but the Subject is as above, with Common Name = Odradek.


The complete server is at tlsserver.c, along with the certificate and key files appcert.pem and appkey.pem (both stored with an additional .text suffix to prevent an accidental click from loading them directly into a browser). To compiler the server under Linux, with the OpenSSL library installed in the usual place, we use

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gcc -o tlsserver tlsserver.c -lssl -lcrypto

The server should be run in the same directory with its certificate and key files. In the following we go through the important lines of the full program stripped of any error checking. Most OpenSSL library methods return 1 on success and 0 on failure. Different kinds of failure may require different error-message libraries.


This does what it sounds like: initializes the SSL subsystem. Loading error strings may also be useful. The next steps commit us to the versions of SSL we agree to accept.

method = SSLv23_server_method();
ctx = SSL_CTX_new(method);
SSL_CTX_set_options(ctx, SSL_OP_NO_SSLv2 | SSL_OP_NO_SSLv3 | SSL_OP_NO_TLSv1 ãÑ| SSL_OP_NO_TLSv1_1);


Somewhat paradoxically, SSLv23_server_method() accepts all SSL and TLS versions. In the third line above, we then disable everything earlier than TLS version 1.1. In openssl version 1.1.0 (the numbering is unrelated to the TLS versioning), SSLv23_server_method() can be replaced with the more appropriately named TLS_server_method().

The variable ctx represents a TLS context, which is a set of TLS state information. Our server will use the same context for all incoming connections.

Next we load the server certificate and key files into our newly created TLS context. Recall that the server side gets our application certificate appcert.pem; the client side will get our certificate-authority certificate CAcert.pem.

int cfile_result = SSL_CTX_use_certificate_file(ctx, "./appcert.pem", SSL_ ãÑFILETYPE_PEM);
 int kfile_result = SSL_CTX_use_PrivateKey_file (ctx, "./appkey.pem", SSL_ ãÑFILETYPE_PEM);


The server needs the key file to sign messages. Next we create an ordinary TCP socket listening on port 4433; createSocket() is defined at the end of the tlsserver.c file.

sock = createSocket(4433);


Finally we get to the main loop. We accept a TCP connection, create a new connection-specific SSL object from our context, and tie the new SSL object to the socket.

while(1) {
int childsock = accept(sock, (struct sockaddr*)&addr, &len);
SSL* ssl = SSL_new(ctx);
SSL_set_fd(ssl, childsock);
SSL_write(ssl, reply, strlen(reply)); }


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SSL_accept() is where the handshaking described in TLS Connection Setup takes place. At this point, the server writes its reply message over the now-encrypted channel, and is done.

We argued in 1.15 IETF and OSI that TLS can in some ways be regarded as a true Presentation layer. Note, however, that the Application layer here (that is, tlsserver.c) is responsible for accepting childsock, and passing it to TLS through SSL_set_fd(); the Application layer, in other words, does interact directly with TCP.


The complete client is at tlsclient.c; its certificate is CAcert.pem (again with a .text suffix). Again we go through the sequence of SSL library calls with error-checking removed. As with the server, we start with initialization; this time, we also load error strings:

SSL_library_init(); // "SSL_library_init() always returns '1' ERR_load_crypto_strings();

Next we again choose what TLS versions we will allow, this time starting with SSLv23_client_method():

method = SSLv23_client_method();
ctx = SSL_CTX_new(method);
SSL_CTX_set_options(ctx, SSL_OP_NO_SSLv2 | SSL_OP_NO_SSLv3|SSL_OP_NO_ ãÑ


If we instead use method = TLSv1_1_client_method(), the connection should fail, this call allows only TLS version 1.1, and the server requires TLS version 1.2 or better.

The next step is to load the trust store, that is, the certificates from the certificate authorities we have elected to trust. If we do nothing, the trust store will be empty. We first load a standard certificate directory (directories are supplied to SSL_CTX_load_verify_locations() as the third parameter and individual files as the second). Certificates in a directory must be named (possibly using symbolic links) by their hash values; see the c_rehash utility. If all we wanted was to be able to trust our own server’s appcert.pem, we could just load our own certificate-authority certificate CAcert.pem, but we will need the standard certificate directory if we want to point tlsclient at a real HTTPS server rather than at tlsserver.

SSL_CTX_load_verify_locations(ctx, NULL, "/etc/ssl/certs")


Next we load our own certificate CAcert.pem, which is then added to the trust store, in addition to the standard certificates. We can add multiple individual certificates by making multiple calls, or by concatenating all the certificates into a single file. The certificates must be separated by their BEGIN CERTIFICATE and END CERTIFICATE lines.

SSL_CTX_load_verify_locations(ctx, "CAcert.pem", NULL);


Now we’re ready to connect to the server. The last line below initiates the TLS handshake, starting with ClientHello.

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sock = openConnection(hostname, port);
ssl = SSL_new(ctx);
SSL_set_fd(ssl, sock);

Next the client retrieves (and, in our case prints) the certificate supplied by the server. If the server is tlsserver, this will normally be appcert.pem, with CN=Odradek.

cert = SSL_get_peer_certificate(ssl);
certname = X509_get_subject_name(cert);
 // print certname


The printed certname does indeed show CN=Odradek, from appcert.pem. If we were writing a web browser, this is the point where we would verify that the site hostname matches the CN field of the certificate. After the client receives the application certificate,

it must verify its signature, with the call below. This is where the client uses its trust store.

ret = SSL_get_verify_result(ssl);


If one of the certificate authorities in the trust store vouches for the signature on the application certificate, the return value above is X509_V_OK, and all is well. If we comment out the loading of CAcert.pem, however, we get “unable to get local issuer certificate”. If, with tlsclient still not loading CAcert.pem, we have the server send CAcert.pem (and CAkey.pem), instead of its proper certificate, we get an error “self-signed certificate”. A full list of certificate-verification errors is listed with the verify command.

If validation fails, the connection is still encrypted, but is vulnerable to a man-in-the-middle attack. Regardless of what happened during validation, our particular tlsclient goes on to write an HTTP GET request to the server, and then read the server’s response (the companion tlsserver program does not in fact read the GET request). Generally speaking, however, continuing the TLS session after a certificate validation failure is a very bad idea.

SSL_write(ssl, request, req_len); // ignored by tlsserver program
do {
bytesread = SSL_read(ssl, buf, sizeof(buf));
fwrite(buf, bytesre
ad, 1, stdout);
  }   while(bytesread > 0);


The written request here is ignored by tlsserver; it is an HTTP GET request of the form GET / HTTP/ 1.1\r\nHost: hostname\r\n\r\n. If we point tlsclient at a real webserver, say

tlsclient 443


then we should again get an X509_V_OK verification result because we loaded the default certificateauthority library.

We can also point the built-in openssl client at tlsserver; by default it connects to localhost at port 4433:

openssl s_client

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Of course, verification fails. This is because s_client doesn’t know about our certificate authority. We can add it, however, on the command line:

openssl s_client -CAfile CAcert.pem

























































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