Comprehensive DNS Resolver Defenses Against Cache Poisoning
draft-weaver-dnsext-comprehensive-resolver-00

Status of this Memo

By submitting this Internet-Draft, each author represents that any applicable patent or other IPR claims of which he or she is aware have been or will be disclosed, and any of which he or she becomes aware will be disclosed, in accordance with Section 6 of BCP 79.

Internet-Drafts are working documents of the Internet Engineering Task Force (IETF), its areas, and its working groups. Note that other groups may also distribute working documents as Internet-Drafts.

Internet-Drafts are draft documents valid for a maximum of six months and may be updated, replaced, or obsoleted by other documents at any time. It is inappropriate to use Internet-Drafts as reference material or to cite them other than as “work in progress.”

The list of current Internet-Drafts can be accessed at http://www.ietf.org/ietf/1id-abstracts.txt.

The list of Internet-Draft Shadow Directories can be accessed at http://www.ietf.org/shadow.html.

This Internet-Draft will expire on April 3, 2009.

Abstract

DNS resolvers are vulnerable to many attacks on their network communication, ranging from blind attacks to full men-in-the-middle. Although a full man-in-the-middle can only be countered with cryptography, there are many layers of defenses which apply to less powerful attackers. Of particular interest are defenses which only require changing the DNS resolvers, not the authoritative servers or the DNS protocols. This document begins with a taxonomy of attacker capabilities and desires, and then discusses defenses against classes of attackers, including detecting non-disruptive attacks, entropy budgeting, detecting entropy stripping, semantics of duplication, and cache policies to eliminate "race-until-win" conditions. Proposed defenses were evaluated with traces of network behavior.

Table of Contents

1.  Introduction
2.  A Taxonomy of Attacks
    2.1.  Passive or Active Attacks
    2.2.  Blind or Aware Attacks
    2.3.  Non-Disruptive or Disruptive Attacks
    2.4.  Transaction or Cache Attacks
    2.5.  Direct Mapping or Ancillary Data Attacks
3.  Race-Until-Win Blind Attacks
4.  Requirements for Blind Transaction Attacks
5.  General Evaluation Strategy
6.  Directly Detecting Non-Disruptive Attacks
7.  Entropy Budgeting
8.  Entropy Stripping
9.  On Duplication for Entropy Increase
10.  Cache Policy, Scoping, and Ancillary Data Attacks
    10.1.  Preliminary Estimates of Performance Impact
    10.2.  Optimization: Only One A Record for the NS RRSet
    10.3.  Optimization: Object Scope
    10.4.  Optimization: Lazily Fetching the NS RRSet
    10.5.  Accepting Requests for Change
11.  Conclusions
12.  Acknowledgements
13.  IANA Considerations
14.  Security Considerations
15.  References
    15.1.  Normative References
    15.2.  Informative References
§  Author's Address
§  Intellectual Property and Copyright Statements

1.  Introduction

DNS resolvers are susceptible to many attacks on their network traffic, ranging from an attacker performing blind packet injection to a full man-in-the-middle, capable of controlling all traffic the resolver receives.

With the recent discovery of the Kaminski Attack (US-CERT, “Multiple DNS implementations vulnerable to cache poisoning,” July 2008.) [kaminski], new attention has been focused on securing DNS from adversaries. This document focuses on a subset of the problem: securing DNS resolvers without changing the DNS authoritative servers or protocols, including authorities that do not actively follow the DNS specification.

This document begins with a taxonomy of attacker properties (Section 2 (A Taxonomy of Attacks)), observations on race-until-win blind attacks (Section 3 (Race-Until-Win Blind Attacks)), the limitations of blind transaction attacks (Section 4 (Requirements for Blind Transaction Attacks)), the evaluation strategy used to study possible defenses (Section 5 (General Evaluation Strategy)), directly detecting non-disruptive attacks (Section 6 (Directly Detecting Non-Disruptive Attacks)), entropy budgeting (Section 7 (Entropy Budgeting)), detecting entropy stripping (Section 8 (Entropy Stripping)), the effects of duplication on protecting the transaction and cache while maintaining compatibility (Section 9 (On Duplication for Entropy Increase)), and finally, the effects of cache policies which resist race-until-win attacks (Section 10 (Cache Policy, Scoping, and Ancillary Data Attacks)).

The key words "MUST", "MUST NOT", "REQUIRED", "SHALL", "SHALL NOT", "SHOULD", "SHOULD NOT", "RECOMMENDED", "MAY", and "OPTIONAL" in this document are to be interpreted as described in RFC 2119 (Bradner, S., “Key words for use in RFCs to Indicate Requirement Levels,” March 1997.) [RFC2119].

2.  A Taxonomy of Attacks

Not all attacker capabilities are equal, and different defenses are able to block some classes of attackers. By forming a taxonomy, one can describe the classes of attackers which each defense can detect or mitigate.

2.1.  Passive or Active Attacks

A passive attacker must wait for the targeted resolver to make requests, while an active attacker is able to trigger specific requests by the resolver. In general, there are numerous mechanisms which an attacker can use to trigger requests. One must assume that, should the attacker desire it, the attacker can be active.

2.2.  Blind or Aware Attacks

A blind attacker does not know any of the entropy sources (e.g. the Transaction ID, port selection, capitalization) of the request. An aware attacker knows this information because he can directly observe the request of the resolver. Increasing the entropy of a request increases the difficulty for blind attackers, but has no effect on aware attackers.

2.3.  Non-Disruptive or Disruptive Attacks

A non-disruptive attacker is unable to block either the resolver's request or the authority's response, while a disruptive attacker can halt either the legitimate request or the legitimate response. Disruptions can take the shape of a DOS attack (if the attacker is not on the packet path), taking advantage of a failure of the authoritative DNS server, mechanisms in the physical layer which enable sequelching a sender, or dropping packets if the attacker is a man-in-the-middle.

2.4.  Transaction or Cache Attacks

A transaction attack targets the individual transaction requested by the end-client: the attacker needs to get his response immediately accepted by the victim application, while a cache attack requires that the resolver cache the attacker's result for future use.

Transaction attacks are often less powerful than cache attacks: A transaction attack targets the single victim request, while cache attacks can have lasting effects until the TTL expires. For blind attacks, as discussed in Section 4 (Requirements for Blind Transaction Attacks), blind transaction attacks are strictly less powerful than blind cache attacks.

2.5.  Direct Mapping or Ancillary Data Attacks

A direct attack targets the immediate answer to the question asked by the resolver. An ancillary data attack requires that the resolver accept some data, be it an NS RRSet, a CNAME, an A record, or other result, which is not the direct answer to the question. A transaction attack must target the direct mapping, while cache attacks can target ancillary data.

3.  Race-Until-Win Blind Attacks

The Kaminski attack is a blind, cache attack targeting ancilllatory data. The additional power of the Kaminski attack is not a reduction in the number of packets required for a successful blind attack, but a reduction in time. Rather than only running one race-condition attack every TTL, the Kaminski attack and variations all rely on ancillary information to achieve a "race-until-win" property: when the attack fails, it can be immediately retried without delay instead of waiting for the TTL to expire.

If a blind attacker targets an actual record, that race can only be run once every TTL, as subsequent requests will simply hit in the cache. Thus, creating any race-until-win attack on a single name requires targeting ancillary data.

4.  Requirements for Blind Transaction Attacks

A blind transaction attack is strictly harder than a blind cache attack. In a blind cache attack, the attacker only needs to coerce the victim into generating DNS requests. A blind transaction attack not only requires coercing the victim into generating a DNS request, but also requires that the victim act upon the result of that request.

Blind transaction attacks are also less powerful than blind cache attacks. In order to target a name with a race-until-win attack, the attacker must be able to not only get the victim to generate a DNS request, and act on that request, but that request must be valid for any arbitrary subname within the target domain. If the blind transaction attack is targeting a single name, it can never be run with a race-until-win property, as it must target the direct mapping and not ancillatory data.

However, there does exist at least one significant blind transaction attack which can be conducted with a "race-until-win" property: targeting cookies contained in a web browser. In this attack, the attacker coerces the victim into visiting a site of the attacker's choosing. This site opens up an iFrame which points to 1.www.target.com, which the attacker attempts to poison with a blind attack. If it fails, the Javascript on the site creates a second iFrame, 2.www.target.com. This process iterates until success, whereupon the victim's browser contacts the attacker's web site, presenting all relevant cookies. Since each attempt uses a different name, the attacker can try continuously until successful.

Although such cookie stealing is noteworthy, any site which allows cookies to be stolen in this manner is also trivially vulnerable to many attack tools such as Cookie Monster (Perry, M., “Fully Automated Active HTTPS Cookie Hijacking,” August 2008.) [perry], and any web service which resists such tools also resists this attack. It is unclear whether the DNS infrastructure should be concerned with this particular attack.

It is also unclear what other blind transaction attacks are possible with this "race-until-win" property, but it relies on the end-application trusting arbitrary subnames and subdomains for attacker-triggered requests. Without this property, the attacker cannot perform a "race-until-win" blind transaction attack. Thus, blind transaction attacks, although they need to be considered, are a far narrower threat than blind cache attacks.

5.  General Evaluation Strategy

We evaluate multiple defenses in this document using network traces analyzed by "Bro", a network analysis environment primarily used for intrusion detection. Bro includes a detailed analyzer for DNS behavior, coupled to a scriptable analysis language.

The primary evaluation was conducted on traces of the ICSI border, capturing 18,329,249 UDP external DNS requests generated by ICSI's resolvers over a period of 19 days. This initial evaluation ignored PTR lookups.

6.  Directly Detecting Non-Disruptive Attacks

Any non-disruptive attack will, with high probability, have the resolver receive two separate and valid responses: one from the attacker and one from the authoritative server. This will occur on blind attacks where the attacker does not suppress the correct operation of the authoritative resolver, such as by flooding it to achieve a denial-of-service. It will also occur if an attacker is acting as a packet injector on a broadcast media if the attacker is not able to squelch the sender.

The presence of a changed duplicate response, where the second response is different from the first response within a short period of time can thus be used to directly detect that a non-disruptive attack has occurred on a transaction, and mitigate any poisoning of the cache.

The algorithm is simple: the resolver maintains a list of all answered responses for a timeout which should be on the order of one or two seconds. If a second response for a request arrives whose transaction and other entropy matches the accepted response but whose value is different, this should be treated as an attack, and all cache entries set or dependent on the first result should be voided, mitigating the attack's effect on the cache. Thus a transaction attack is detected but not mitigated (as the final victim may have acted on the result), but a cache attack is both detected and mitigated

Furthermore, if the resolver is a recursive resolver, it should also forward the second response to the originator with the TTL changed to 0 seconds, which allows the initial questioner, if running the same algorithm, to also know that the result was compromised, enabling the end system to detect the attack on the transaction and mitigate any effects on a local cache.

We developed a Bro analysis script to detect this condition to evaluate the false positive rate. We observed only one such site at ICSI, a realtime DNS blackhole service. A run at Ohio State did found a few other sites, which we haven't yet assessed.

As importantly, if resolvers implemented this approach, they would fail to cache only a handful of authorities which present this anomaly on normal results. Until the stub resolvers are updated to treat this condition as an attack, these sites would still resolve successfully but would not be cached.

The state-holding requirements, although nontrivial, are also reasonable. Consider a recursive resolver (or a cluster node in a clustered resolver implementation) that generates 10,000 outstanding external queries per second. If the timeout is 1 second, and the state-holding required involves 1 KB of data, the additional memory requirements would only be 10 MB of data for such a resolver. Even with a longer timeout or a more active resolver, the state-holding requirements should be reasonable.

Furthermore, although the replies themselves need to be maintained, for the most part this information is already retained in the cache. For any data element that is cached, only a pointer to that element needs to be stored to perform this consistency check, rather than the full answer.

This should also not affect any resolver composed of a cluster of systems as long as the load balancer is deterministic: always sending the second response back to the cluster node which received the first response. This enables each cluster node to maintain its own list of replies to check for changed responses.

This policy would have no effect on the latency of resolvers. Until end-user stub resolvers and applications are updated to treat this as an attack, even sites with anomalous DNS authorities will still resolve properly, but may experience slightly higher load due to lack of cacheability.

7.  Entropy Budgeting

An important question is "how much entropy is necessary" to prevent blind attacks. It's obvious that 16 bits (16b) of entropy in a query is insufficient, both to protect transactions and to protect the cache. 32b of entropy may in practice currently suffice to mitigate most attacks today, as attackers will prefer victims that do not implement proper randomization.

Thus the question, assuming all resolvers implement increased entropy defenses, what is sufficient query entropy to protect the cache? We argue that 40b of entropy is more than sufficient. As the probability of attack success with K tries and N bits of entropy is 1-(1-1/2^N)^K, 40b of entropy requires nearly a terapacket of attacks to be successful with reasonable probability.

40b of entropy is easily obtainable for resolvers not behind entropy-stripping devices such as NATs (Section 8 (Entropy Stripping)). With 16b of entropy for the transaction ID, over 15b of entropy for source-port randomization, and names of at least 9 characters using 0x20 randomization (observing that most DNS authoritative servers maintain capitalization, thus random capitolization can provide a nonce value)[0x20] (Vixie, P. and D. Dagon, “Use of Bit 0x20 in DNS Labels to Improve Transaction Identity,” March 2008.) achieves the necessary 40b threshold without resorting to duplication in most cases (Section 9 (On Duplication for Entropy Increase)).

It is possible for a resolver to use a simple count of failures, responses which match requests and appear to come from the proper authority but have different entropy values, to know that it is under attack and respond with duplication Section 9 (On Duplication for Entropy Increase) while the attack is ongoing, well before an attacker could have an effect on the cache with large entropy. In our observations on ICSI traces we noticed that some authoritative servers do this naturally, but the rate is low.

If the resolver responded to 1k PPS of attacks with duplication, and the entropy budget is 40b, an attacker attempting to go below the threshold by sending 990 attacks per second would need over 2 hours to have even a .001% chance of success, or over 80 days of a continual attack to have just a 1% chance. Yet by requiring the attacker to send 1k packets-per-second to trigger duplication in the resolver, duplication which only needs to be in place while under attack, this shouldn't prove to be an effective DOS.

However, the author initially advocated clearing the cache at a threshold of attack. This was an error, as voiding the cache does not provide a benefit as the attacker knows when his attack is successful, and could accept the voided cache and just keep trying until successful.

8.  Entropy Stripping

Many network devices, such as NATs, firewalls, and mandatory proxies, can exist between a resolver and the remote authority it is querying. Some of these devices will rewrite important information, such as IP addresses, UDP port numbers, or even DNS transaction IDs, fields that the resolver relies on as sources of query entropy. These devices are likely to be located near one or the other endpoint system.

If a DNS resolver is attempting to increase its request entropy by using one or more of these sources, the resolver must know if these entropy sources are being stripped from its network communication.

The simple solution is to query a specialized authority which returns the entropy value it receives as an A record. An example of such an authority has been temporarily operating at {port,id,server}.nettest.icir.org, where the transaction ID and port are returned as the least significant two bytes of the address, while server returns the IP address of the contacting DNS resolver. An alternate version, entropy.nettest.icir.org, returns a CNAME to a human readable form of all three values, while respecting the incoming capitalization to verify 0x20 operation.

When a DNS resolver starts, and when a resolver notices that its IP address has changed, it should query such an authority, provided by either the resolver's software developer or a third party, to determine if it is located behind a NAT or other entropy-stripping device. If the resolver is behind a NAT, it must then use duplication (Section 9 (On Duplication for Entropy Increase)) to protect the cache if the remaining entropy is not sufficient to meet the security goals (Section 7 (Entropy Budgeting)).

A case where this does not provide protection is if the authority and the attacker, not the resolver, is behind an entropy-stripping network device. In such a case, an attacker capable of forging packets from within the authority's network is likely able to perform other activities far more damaging than a blind attack on DNS requests from that authority. Also, the entropy stripping is only affecting queries to this authority, not all queries performed by a resolver. This also assumes than any entropy-stripper is not malicious and would therefore not benefit from actively whitelisting the test.

9.  On Duplication for Entropy Increase

If 40b of entropy are available on the request and the resolver is not under a significant attack, duplication is not necessary. Thus duplication should be viewed as a fallback position for resolvers which are behind a NAT or other entropy stripping device, accessing authoritative servers which don't respond to 0x20 (per Dagon et. al's proposal [0x20] (Vixie, P. and D. Dagon, “Use of Bit 0x20 in DNS Labels to Improve Transaction Identity,” March 2008.), or is known to be under an active blind attack.

Most authorities are deterministic to multiple queries: if two requests for the same name are received in a short period of time, the returned values will be identical. By issuing two identical requests with different entropy values, this nearly doubles the entropy if we require that both responses are identical before acceptance. Specifically, if the request has K bits of entropy, two requests which are accepted if the responses are identical has 2K-1 bits of entropy.

Dagon et. al.'s 0x20 proposal [0x20] (Vixie, P. and D. Dagon, “Use of Bit 0x20 in DNS Labels to Improve Transaction Identity,” March 2008.) uses this to increase query entropy when capitalization is not preserved by an authority, as without 0x20 the available entropy is only around 32b per query. Similarly, duplication can be used in any case where there is insufficient entropy, such as the impact of an entropy-stripping network device, or if the resolver knows it is currently under a blind attack.

Unfortunately, not all authorities are deterministic in this manner, including some critical authoritative servers belonging to Content Distrbution Networks such as Akamai[akamai] (Akamai Inc, “The Akamai CDN,” 2008.) and others that use frequently changing DNS responses for load-balancing. If two distinct responses are received, and the resolver randomly selects one, this reduces by 1 bit the entropy of the request when compared with no duplication. Thus, for purposes of the cache, it is clear that a resolver must not cache a response when the two responses are different.

However, it appears reasonable to return one of the values as the answer to the transaction, as blind transaction attacks are less powerful than blind cache attacks (Section 4 (Requirements for Blind Transaction Attacks)). For a blind transaction attack to work, the attacker must target a domain served by such a volatile authority as well as coerce the victim to act on this. Although this does leave a small window of vulnerability open, it is proabably preferable to the alternative of not resolving thees names.

By setting the TTL to 0 and returning a randomly selected response, this enables compatibility with nondeterministic authorities without compromising the integrity of either the resolver's cache or the final client's cache. Since it also only reduces the transactional entropy by 1b, it does not make transactions significantly less secure than without duplication.

The alternate approach, iterate until convergence as proposed by Barwood (Barwood, G., “The Birthday Defense (IETF NameDroppers Mailing List),” September 2008.) [barwood], will succeed with very high probability if the remote resolver returns single answers, but fails when the authority returns RRSets which contain more than one element where the RRSets are randomly chosen, if a complete match on the RRSet is required. If "return one of N of the RRSets" is employed it works well.

10.  Cache Policy, Scoping, and Ancillary Data Attacks

RFC 2181 section 5.4.1 (Elz, R. and R. Bush, “Clarifications to the DNS Specification,” July 1997.) [RFC2181] is the current specification for how to accept ancillatory data. It allows ancillatory data, if "in bailywick" (from a server which should be an authority for this name), to be cached for subsequent use, although it should not be returned as an answer. It is this caching of ancillatory data that enables the "race-until-win" Kaminski blind cache attack. Likewise, before bailywick-checking was deployed, ancillatory data was used for classic glue poisoning.

The ancillatory data includes all the data beyond the direct answer to the query (including the NS RRSet, A records associated with the NS RRSet, A records associated with a CNAME for the direct answer, or any other data). This data serves four purposes:

• The NS RRSet and associated A records needed to resolve the current request.
• Items, such as an A record for a CNAME alias, which if accepted will speed the current request's processing by removing the need to fetch additional records.
• Items which if placed in the cache will speed subsequent lookups.
• An indication that an item in the cache is now obsolete.

This suggests that items have a different role in two scopes, analogous to how programming languages view scope: the local scope of the current recursive query and the global scope of the cache. Within the local scope, for purposes of resolving the direct question, ancillary data often must be trusted: otherwise authoritative nameservers may not be reachable when they exist within the domain being queried or in other cases where domains host each-others nameservers. Yet for the purposes of resolving a query, if an authority lies about the ancillatory data, it could just as easily lie about the direct answer, making this data no less trustworthy for processing this answer.

Yet for purposes of inclusion into the global scope, or for returning as the response to a query, ancillary data must not be trusted. It is, by definition, unsolicited information and not an authoritative response, and lies at the heart of the Kaminski attack. If a recursive resolver never accepts ancillary data into the cache, it becomes impossible to target a single name with a race-until-win blind attack.

However, a resolver may safely perform an independent fetch for any piece of ancillary data. This second fetch must not reuse the local scope of the previous fetch, but instead be fetched using a new local context. This enables both the use of ancillary data in responses (such as A records involved in CNAMES) and to speed subsequent responses (such as obtaining the NS RRSet for subsequent lookups).

As an example, suppose the query for 'www.example.com' returns 'www.example.com CNAME server.example.com, example.com NS ns.example.com, ns.example.com A 12.34.56.78, server.example.com A 12.34.56.90'. The resolver can succesfully cache and return 'www.example.com CNAME server.example.com', but must perform independent queries to fetch the NS RRSet and the A records for ns.example.com and server.example.com, and it can't return 'www.example.com CNAME server.example.com, server.example.com A 12.34.56.90' to the final client until the lookup for server.example.com completes.

As a result, all information which is either placed in the global scope or returned to the final client will be validated directly by querying the proper authoritative server. There are no race-until-win attacks possible for including a name into the cache, as inserting a name only would take place if it doesn't exist in the cache which limits any attempt to once every TTL.

This is more restrictive than the policy in RFC 2181 section 5.4.1 (Elz, R. and R. Bush, “Clarifications to the DNS Specification,” July 1997.) [RFC2181]. The RFC policy allows ancillary data to be accepted into the global scope (cache) for purposes of subsequent query processing, which enables race-until-win attacks.

Although the RFC states that the final client should also not accept these authority or additional records, it is unclear whether the stub resolver follow the RFC. Thus a well structured recursive resolver should return results which are safe even if the client does cache additional records in violation of the RFC.

The simplest mechanism for validation is to perform an explicit fetch, a policy being implemented in Unbound (Wijngaards, W., “Resolver Side Mitigations,” August 2008.) [unbound]: For CNAMEs and A records that are not the direct answer to the query, such records are not accepted directly but are instead fetched independently. For the NS record and associated A records, the separate global and local scope is approximated by validating the NS RRSet and all elements within it using subsequent separate fetches.

There is some minor uncertanty if this is a close-enough approximation for the NS RRSet. There exists a short-term time window where the NS RRSet and A records would be in the cache but unvalidated, which may provide a narrow opportunity for abuse. However, this policy does not require new data structures to maintain the local scope of a recursive query, and the abuse window is probably sufficiently small.

Note that if an explicit local scope is maintained, such a policy of fetching all ancillary data rather than including it subsumes bailiwick checking for accepting data into the cache, as no data is included into the cache for any use that was not explicitly requested from an authoritative server.

10.1.  Preliminary Estimates of Performance Impact

Such a policy, although protecting from race-until-win conditions, does impose a higher load on the DNS infrastructure. For queries whose direct answer is not a CNAME, it does not add any latency to query processing if the resolver returns an empty authoritative NS RRset when the NS RRset has yet to be validated, but the number of outstanding queries is significantly increased.

We developed a Bro analysis script to estimate the impact of this policy. This policy accepted a trace of DNS requests and replies and estimated the number of additional queries that would be generated to follow CNAME chains, to cache the NS RRSet and associated A records, and to cache any A records not associated with either a CNAME chain or the NS RRSet. It creates a model of what the resolver's own cache looks like and uses this to estimate the load from additional requests.

For ICSI's 18,329,249 queries, this simple policy would require an additional 16% additional fetches for NS records, an additional 10% of fetches for A records associated with the NS RRSets, .7% additional fetches of A records not associated with NS RRSets and a trivial number of CNAMEs. Thus this policy increases the number of requests by 27%, a nontrivial but still modest overhead. Later on, we discuss some policies that can significantly reduce the number of fetches while maintaining safety.

Only a few queries would have their latency increased if the resolver does not return the NS RRset to the client for the first query because it is still unvalidated. Thus the only cases where a query would see increased latency is when the answer is a CNAME and the record pointed to by this CNAME is in the record. Only .2% of ICSI's queries showed this behavior.

If the resolver does include the NS RRSet it needs to wait for it to be validated in case the final client caches the names for other purposes. This would increase the latency on 16% of the queries which require accessing an external authority to obtain the NS RRSet.

10.2.  Optimization: Only One A Record for the NS RRSet

It is obviously excessive to fetch all A records for the NS RRSet: unless there is a failure, only one nameserver needs to be contacted to return future results, although that resolver might be sub-optimal from a latency viewpoint. If only one A record is fetched for the NS RRSet rather than all A records, this reduces from 10% to 6% the number of additional records fetched for the purpose of providing a valid name associated with each NS RRset.

However, this may increase the latency of subsequent responses if the chosen authoritative server is not the closest. This latency could be mitigated by performing additional fetches of alternate nameservers as requests for a domain continue to be made. One suggestion is to use the "closest name", to ranodmly select the name which most closely matches the domain. Thus if "example.com NS ns.example.com, ns.example2.com, ns.example3.org", needed to be cached, ns.example.com would be fetched. If this lookup failed, ns.example2.com would be fetched, followed by ns.example3.org: selecting from the most matching to least-matching name. It is unclear what the latency penalty would be for this heuristic.

10.3.  Optimization: Object Scope

There does exist an additional data scope: object local scope, that can act as an optimization. For example, the A records associated with an NS RRSet cannot be accepted directly. Yet since they were returned in association with a specific query for the NS RRSet, they can be trusted solely for purposes of evaluating the NS RRset.

As an example, if the response for a query for the nameservers of example.com is 'example.com NS ns.example.com, ns.example.com A 12.34.56.78', it is acceptable to use the A record solely for the purposes of a subsequent lookup of a value in example.com, but not for any other purposes. This, in programming language terms, represents object-local scope: a data value that can only be used in the context of another data value.

This arrises from a simple observation: if the response is corrupted, any value in the response could have been corrupted, including the legitimate answer. Thus there is no risk in using the ancillatory data in a context where the direct answer was trusted. But ancillatory data can not be trusted outside the context where the direct answer is trusted, because this enables race-until-win attacks.

There are two mechanisms where support for object-local scope can enhance resolver performance without compromising safety. The first is to collapse CNAMES and return the result, per Ohta's suggestion (cite). In this, an authority that returns "www.example.com CNAME a.example2.com, a.example2.com A 10.34.56.78" could be treated by the resolver (and returned to the final client) as "www.example.com A 10.34.45.78". This acts to eliminate the latency penalty for fetching CNAME chains.

The second area where object-local scope provides significant savings is for the NS RRset. When the NS RRset is requested, the A records can't be considered as authoritative in general, but MAY be considered as valid only for the use of the NS RRset for subsequent name lookups. This completely eliminates the need to look up the A records associated with the NS RRSet, while still preventing auxilary data from poisoning the cache. However, it does require changes to cache architectures to support this notion: the NS RRset records must include inernal A records which are not exported to the rest of the cache.

The savings are significant: support for object-local scope allows the resolver to not fetch the A records for the NS entries, reducing the total overhead from 27% to 17% for ICSI's traffic

10.4.  Optimization: Lazily Fetching the NS RRSet

There exists a latency/performance tradeoff in fetching the NS RRset. Observe that, for many sites, only a single name is used. Thus, fetching (and caching) the NS RRset represents wasted effort. Instead of eagerly fetching the NS RRset, fetching the NS RRset and any associated A records can be delayed until a second query is generated.

The traffic savings are substantial. In the ICSI traces, instead of requiring 16% additional fetches of the NS RRset, only 4% additional fetches would have been required, with a similar reduction in the number of A records which need to be fetched. As a consequence, however, the lookup of the second name will be delayed as the NS RRset needs to be fetched first. Yet the penalty is fairly modest, as only the second (and not subsequent) fetches would incur this latency penalty. Thus only 4% of the queries would see this latency penalty.

10.5.  Accepting Requests for Change

The final use of ancillary records is to indicate a change request: a statement that a previously cached value is no longer valid despite still being within the TTL of the item. The use of ancillatory data to indicate changes is somewhat out of the protocol specification, but is often considered essential behavior.

It is clear that a resolver must not overwrite an item in the cache with an ancillary item for any reason, as otherwise ancillary data can be used for a race-until-win attack using data replacement instead of data inclusion.

It is also clear that a resolver must not void a cache entry upon a change request for an ancillary item when the response is not in bailiwick. Otherwise, an attacker who controls an arbitrary authority could construct a race-until-win attack by alternating between attacking the name and performing an unrelated query which uses the attacker's authority uses to void the name.

However, a resolver may (and probably should) safely respond to an in-bailiwick request for change by voiding the cache entry associated with the ancillary item. A blind attacker cannot use this behavior to create a race-until-win condition, as the attacker would have to win a race against an arbitrary name in the same domain to void the cache entry and then win the next subsequent race to set the cache entry to the attacker's desired value. As winning two back-to-back races is exponentially harder than winning a single race, the policy is safe from attack while still enabling ancillary data to act as a notification of change.

11.  Conclusions

There are many defenses which can be layered to provide robust defenses for recursive DNS resolvers.

By looking for duplicate responses to the same transaction which have different value, all non-disruptive attacks can be directly detected, and their effects on the cache mitigated. By forwarding the second response, the final victim can also be notified of the attack. This requires ~10MB of state for a resolver performing 10,000 external queries a second.

By setting an entropy budget of 40b, blind attacks are infeasible, requiring terapackets to have a high probability. By querying special authoritative DNS servers, a resolver can detect any process which reduces this entropy, and can use duplicate requests to restore entropy.

For the few sites where duplication produces different results, it is probably safe to return a randomly selected result. Although this still enables transactional cache attacks, it is probably better to accept this very narrow window of vulnerability to enable resolution of key sites.

Finally, cache policy can eliminate 'race-until-win' attacks while subsuming most bailywick checks. By only accepting and returning the direct answer to requests, an attacker can no longer conduct any race-until-win attacks targeting specific names. The overhead is reasonable as even without optimizations, a study of our resolvers showed a 26% increase in requestes before any optimizations are applied.

These changes all involve only resolver behavior, and they also combine to provide better protection than any one defense alone. Increasing entropy increases the blind attacker's work in packets, while eliminating race-until-win increases the work in time. With a sufficient entropy budget, a resolver can detect that it is under attack and act according. While directly detecting non-disruptive attacks can detect both packet injectors and many blind injectors.

12.  Acknowledgements

This work is sponsored in part by NSF Grants ITR/ANI-0205519 and NSF-0433702. All opinions are those of the author, not the funding institution.

Feedback from Vern Paxson, Robin Sommer, Christian Kreibich, Paul Vixie (who also suggested the "closest name" heuristic), Seth Hall, Wooter Wijngaards, Dan Kaminski, and David Dagon

13.  IANA Considerations

None

14.  Security Considerations

This text is focused on security concerns for DNS resolvers. The security aspects of each defense are discussed as part of each section.

15.  References

15.1. Normative References

15.2. Informative References

Author's Address

Full Copyright Statement

Copyright © The IETF Trust (2008).

This document is subject to the rights, licenses and restrictions contained in BCP 78, and except as set forth therein, the authors retain all their rights.

This document and the information contained herein are provided on an “AS IS” basis and THE CONTRIBUTOR, THE ORGANIZATION HE/SHE REPRESENTS OR IS SPONSORED BY (IF ANY), THE INTERNET SOCIETY, THE IETF TRUST AND THE INTERNET ENGINEERING TASK FORCE DISCLAIM ALL WARRANTIES, EXPRESS OR IMPLIED, INCLUDING BUT NOT LIMITED TO ANY WARRANTY THAT THE USE OF THE INFORMATION HEREIN WILL NOT INFRINGE ANY RIGHTS OR ANY IMPLIED WARRANTIES OF MERCHANTABILITY OR FITNESS FOR A PARTICULAR PURPOSE.

Intellectual Property

The IETF takes no position regarding the validity or scope of any Intellectual Property Rights or other rights that might be claimed to pertain to the implementation or use of the technology described in this document or the extent to which any license under such rights might or might not be available; nor does it represent that it has made any independent effort to identify any such rights. Information on the procedures with respect to rights in RFC documents can be found in BCP 78 and BCP 79.

Copies of IPR disclosures made to the IETF Secretariat and any assurances of licenses to be made available, or the result of an attempt made to obtain a general license or permission for the use of such proprietary rights by implementers or users of this specification can be obtained from the IETF on-line IPR repository at http://www.ietf.org/ipr.

The IETF invites any interested party to bring to its attention any copyrights, patents or patent applications, or other proprietary rights that may cover technology that may be required to implement this standard. Please address the information to the IETF at ietf-ipr@ietf.org.