### David Fifield*, Chang Lan, Rod Hynes, Percy Wegmann, and Vern Paxson
# Blocking-resistant communication through domain fronting


**Abstract:** We describe “domain fronting,” a versatile
censorship circumvention technique that hides the remote endpoint of a communication. Domain fronting
works at the application layer, using HTTPS, to communicate with a forbidden host while appearing to communicate with some other host, permitted by the censor. The key idea is the use of different domain names at
different layers of communication. One domain appears
on the “outside” of an HTTPS request—in the DNS request and TLS Server Name Indication—while another
domain appears on the “inside”—in the HTTP Host
header, invisible to the censor under HTTPS encryption. A censor, unable to distinguish fronted and nonfronted traffic to a domain, must choose between allowing circumvention traffic and blocking the domain entirely, which results in expensive collateral damage. Domain fronting is easy to deploy and use and does not require special cooperation by network intermediaries. We
identify a number of hard-to-block web services, such as
content delivery networks, that support domain-fronted
connections and are useful for censorship circumvention.
Domain fronting, in various forms, is now a circumvention workhorse. We describe several months of deployment experience in the Tor, Lantern, and Psiphon circumvention systems, whose domain-fronting transports
now connect thousands of users daily and transfer many
terabytes per month.

**Keywords: censorship circumvention**

DOI 10.1515/popets-2015-0009
Received 2015-02-15; revised 2015-02-15; accepted 2015-05-15.

***Corresponding Author: David Fifield: University of**
California, Berkeley, E-mail: fifield@eecs.berkeley.edu
**Chang Lan: University of California, Berkeley, E-mail:**
clan@eecs.berkeley.edu
**Rod Hynes: Psiphon Inc, E-mail: r.hynes@psiphon.ca**
**Percy Wegmann: Brave New Software, E-mail:**
ox.to.a.cart@gmail.com
**Vern Paxson: University of California, Berkeley and**
the International Computer Science Institute, E-mail:
vern@berkeley.edu


## 1 Introduction

Censorship is a daily reality for many Internet users.
Workplaces, schools, and governments use technical and
social means to prevent access to information by the network users under their control. In response, those users
employ technical and social means to gain access to the
forbidden information. We have seen an ongoing conflict
between censor and censored, with advances on both
sides, more subtle evasion countered by more powerful
detection.
Circumventors, at a natural disadvantage because
the censor controls the network, have a point working
in their favor: the censor’s distaste for “collateral damage,” incidental overblocking committed in the course of
censorship. Collateral damage is harmful to the censor,
because the overblocked content has economic or social
value, so the censor tries to avoid it. (Any censor not
willing to turn off the Internet completely must derive
_some benefit from allowing access, which overblocking_
harms.) One way to win against censorship is to entangle circumvention traffic with other traffic whose value
exceeds the censor’s tolerance for overblocking.
In this paper we describe “domain fronting,”
a general-purpose circumvention technique based on
HTTPS that hides the true destination of a communication from a censor. Fronting works with many web
services that host multiple domain names behind a frontend server. These include such important infrastructure as content delivery networks (CDNs) and Google’s
panoply of services—a nontrivial fraction of the web.
(Section 4 is a survey of suitable services.) The utility
of domain fronting is not limited to HTTPS communication, nor to accessing only the domains of a specific web
service. It works well as a domain-hiding component of
a larger circumvention system, an HTTPS tunnel to a
general-purpose proxy.
The key idea of domain fronting is the use of different domain names at different layers of communication.
In an HTTPS request, the destination domain name appears in three relevant places: in the DNS query, in the
TLS Server Name Indication (SNI) extension [18, §3],
and in the HTTP Host header [20, §14.23]. Ordinarily, the same domain name appears in all three places.


-----

In a domain-fronted request, however, the DNS query
and SNI carry one name (the “front domain”), while the
HTTP Host header, hidden from the censor by HTTPS
encryption, carries another (the covert, forbidden destination).

**Fig. 1. Domain fronting uses different domain names at different**
layers. At the plaintext layers visible to the censor—the DNS
request and the TLS Server Name Indication—appears the front
domain allowed.example. At the HTTP layer, unreadable to the
censor, is the actual, covert destination forbidden.example.

The censor cannot block on the contents of the DNS
request nor the SNI without collaterally blocking the
front domain. The Host header is invisible to the censor,
but visible to the frontend server receiving the HTTPS
request. The frontend server uses the Host header internally to route the request to its covert destination;
no traffic ever reaches the putative front domain. Domain fronting has many similarities with decoy routing [29, 35, 69, 70]; it may be understood as “decoy
routing at the application layer.” A fuller comparison
with decoy routing appears in Section 3.
This Wget command demonstrates domain fronting
on Google, one of many fronting-capable services. Here,
the HTTPS request has a Host header for maps.google.
com, even though the DNS query and the SNI in the
TLS handshake specify www.google.com. The response
comes from maps.google.com.
```
  $ wget -q -O - https://www.google.com/ \
    --header 'Host: maps.google.com' | \
    grep -o '<title>.*</title>'
  <title>Google Maps</title>

```
A variation is “domainless” fronting, in which there is no
DNS request and no SNI. It appears to the censor that
the user is browsing an HTTPS site by its IP address,
or using a web client that does not support SNI. Domainless fronting can be useful when there is no known
front domain with sufficiently high collateral damage; it
leaves the censor the choice of blocking an entire IP address (or blocking SNI-less connections entirely), rather
than blocking only a single domain. According to our


communication with the International Computer Science Institute’s certificate notary [32], which observes
on the order of 50 million TLS connections daily, 16.5%
of TLS connections in June 2014 lacked SNI, which is
enough to make it difficult for a censor to block SNI-less
TLS outright.
Domain fronting works with CDNs because a CDN’s
frontend server (called an “edge server”), on receiving
a request for a resource not already cached, forwards
the request to the domain found in the Host header
(the “origin server”). (There are other ways CDNs may
work, but this “origin pull” configuration is common.)
The client issues a request that appears to be destined
for an unrelated front domain, which may be any of
the CDN’s domains that resolve to an edge server; this
fronted request is what the censor sees. The edge server
decrypts the request, reads the Host header and forwards the request to the specified origin, which in the
circumvention scenario is a general-purpose proxy. The
origin server, being a proxy, would be blocked by the
censor if accessed directly—fronting hides its address
from the censor.
On services that do not automatically forward requests, it is usually possible to install a trivial “reflector” web application that emulates an origin-pull CDN.
In this case, fronting does not protect the address of
the origin per se; rather it protects the address of the
reflector application, which in turn forwards requests to
the origin. Google App Engine is an example of such
a service: against a censor that blocks the App Engine
domain appspot.com but allows other Google domains,
domain fronting enables access to a reflector running on
appspot.com.
No matter the specifics of particular web services, as
a general rule they do not forward requests to arbitrary
domains—only to domains belonging to one of their customers. In order to deploy a domain-fronting proxy, one
must become a customer of the CDN (or Google, etc.)
and pay for bandwidth. It is the owner of the covert
domain who pays the bandwidth bills, not the owner of
the front domain, which need not have any relation to
the covert domain beyond using the same web service.
The remainder of this paper is devoted to a deep
exploration of domain fronting as we have deployed
it in practice. Section 2 explains our threat model
and assumptions. Section 3 gives general background
on the circumvention problem and outlines its three
grand challenges: address-based blocking, content-based
blocking, and active probing. Domain-fronting systems
are capable of meeting all three challenges, forcing censors to use more expensive, less reliable censorship tech

-----

niques that have heretofore not been seen in practice.
Section 4 is a survey of CDNs and other services that
are usable for fronting; we identify general principles as
well as idiosyncrasies that affect implementation. The
following sections are three case studies of deployment:
Section 5 for Tor, Section 6 for Lantern, and Section 7
for Psiphon. Section 8 sketches domain fronting’s resistance to statistical traffic analysis attacks. Section 9 has
general discussion and Section 10 summarizes.

## 2 Threat model

Our threat model includes four actors: the censor, the
censored client, the intermediate web service, and the
covert destination (a proxy server). Circumvention is
achieved when the client reaches the proxy, because the
proxy grants access to any other destination. The client
and proxy cooperate with each other. The intermediate
web service need not cooperate with either, except to
the extent that it does not collude with the censor.
The censor controls a (generally national) network
and the links into and within it. The censor can inspect
traffic flowing across all links under its control and can
block or allow any packet. The censor can inject and replay traffic, and operate its own clients and servers. The
client lies within the censor’s network, while the intermediate web service and proxy lie outside. The censor
blocks direct communication between the client and the
proxy, but allows HTTPS between the client and at least
one front domain or IP address on the intermediate web
service.
The client, intermediate web service, and destination proxy are uncontrolled by the censor. The censor
does not control a trusted certificate authority: it cannot
man-in-the-middle TLS without being caught by ordinary certificate validation. The client is able to obtain
the necessary circumvention software.

## 3 Background and related work

Broadly speaking, there are three main challenges in
proxy-based circumvention: blocking by content, blocking by address, and active probing. Blocking by content
is based on what you say, blocking by address is based on
_whom you talk to, and active probing means the censor_
_acts as a client. A savvy censor will employ all these_
techniques, and effective circumvention requires countering them all.


A content-blocking censor inspects packets and payloads, looking, for example, for forbidden protocols or
keywords. Content-based blocking is sometimes called
deep packet inspection (DPI). An address-blocking censor forbids all communication with certain addresses,
for example IP addresses and domain names, regardless of the contents of the communication. An activeprobing censor does not limit itself to observation and
manipulation of user-induced traffic only. It sends its
own proxy requests (active probes), either proactively or
on demand in response to observed traffic, and blacklists
any proxies it finds thereby. Active probing is a precise
means of identifying proxy servers, even when a protocol
is hard to detect on the wire—it can regarded as a way
of reducing accidental overblocking. Winter and Lindskog [66] confirmed an earlier discovery of Wilde [64]
that China’s Great Firewall discovers secret Tor bridges
by issuing followup scans after observing a suspected
Tor connection.
There are two general strategies for countering
content-based blocking. The first is to look unlike anything the censor blocks; the second is to look like something the censor allows. Following the first strategy are
the so-called “look-like-nothing” transports whose payloads look like a uniformly random byte stream. Examples of look-like-nothing transports are obfuscatedopenssh [41] and its string of successors: obfs2 [33],
obfs3 [34], ScrambleSuit [67], and obfs4 [4]. They all
work by re-encrypting an underlying stream so that
there remain no plaintext components, not even in the
handshake and key exchange. obfs2, introduced in early
2012 [14], used a fairly weak key exchange that can be
detected passively; it is now deprecated and little used.
obfs3 is Tor’s most-used transport [61] as of May 2015. It
improves on obfs2 with a Diffie–Hellman key exchange,
public keys being encoded so as to be indistinguishable
from random binary strings. ScrambleSuit and obfs4
add resistance to active probing: the server accepts a
TCP connection but does not send a reply until the
client proves knowledge of a secret shared out of band.
ScrambleSuit and obfs4 additionally obscure the traffic
signature of the underlying stream by modifying packet
lengths and timing.
The other strategy against DPI is the steganographic one: look like something the censor allows.
fteproxy [17] uses format-transforming encryption to encode data into strings that match a given regular expression, for example a regular-expression approximation of
HTTP. StegoTorus [63] transforms traffic to look like
a cover protocol using a variety of special-purpose encoders. Code Talker Tunnel (formerly SkypeMorph) [47]


-----

mimics a Skype video call. FreeWave [30] encodes a
stream as an acoustic signal and sends it over VoIP to a
proxy. Dust [65] uses encryption to hide static byte patterns and then shapes statistical features such as packet
lengths and byte frequencies to match specified distributions.
Houmansadr et al. [28] evaluate “parrot” systems
that imitate another protocol and conclude that unobservability by imitation is fundamentally flawed. To
fully mimic a complex and sometimes proprietary protocol like Skype is difficult, because the system must
imitate not only the normal operation of the protocol,
but also its reaction to errors, its typical traffic patterns, and quirks of implementations. Geddes et al. [23]
demonstrate that even non-parrot systems may be vulnerable to attacks that disrupt circumvention while having little effect on ordinary traffic. Their examination
includes VoIP protocols, in which packet loss and duplication are acceptable. The censor may, for example,
strategically drop certain packets in order to disrupt a
covert channel, without much harming ordinary calls.
The challenge of address-based blocking is a difficult one that has inspired various creative circumvention ideas. Tor has long faced the problem of the
blocking of its relays, the addresses of which appear
in a public directory. In response, Tor began to reserve a portion of its relays as secret “bridges” [15]
whose addresses are not publicly known. BridgeDB [9],
the database of secret bridges, carefully distributes addresses so that it is easy to learn a few bridges, but hard
to enumerate all of them. BridgeDB uses CAPTCHAs
and other rate-limiting measures, and over short time
periods, always returns the same bridges to the same
requester, preventing enumeration by simple repeated
queries. BridgeDB is also capable of distributing the
addresses of obfuscated bridges (currently obfs3, obfs4,
ScrambleSuit, and fteproxy), granting IP-blocking resistance to DPI-resistant systems that otherwise lack it.
CensorSpoofer [62] decouples upstream and downstream data channels. The client sends data to a CensorSpoofer proxy over a low-bandwidth covert channel
such as email. The proxy sends data back over a UDP
channel, all the time spoofing its source address so the
packets appear to originate from some other “dummy”
host. The censor has no IP address to block, because
the proxy’s true address never appears on the wire.
Client and server have the challenge of agreeing on a
dependable covert upstream channel that must remain
unblocked, and the client must carry on a believable
UDP conversation with the dummy host—a VoIP call,
for example.


Flash proxy [22] resists address blocking by
conscripting web users as temporary proxies. Each
JavaScript-based proxy lasts only as long as a user stays
on a web page, so the pool of proxies is constantly changing. If one of them is blocked, there is soon another
to replace it. Flash proxy’s approach to address blocking is the opposite of domain fronting’s: where flash
proxy uses many cheap, disposable, individually blockable proxies, domain fronting uses just a few high-value
front domains on hard-to-block network infrastructure.
A drawback with flash proxy’s use of the browser is
that the client must be able to receive a TCP connection; in particular it must not be behind network address
translation (NAT), which limits flash proxy’s usefulness.
Part of the flash proxy protocol requires the client to
send a small amount of unblockable data in a process
called rendezvous. The default rendezvous mechanism
has used domain fronting through Google App Engine
since 2013 [58]. Flash proxy itself does nothing to defend
against DPI. Connections between censored clients and
browser-based proxies use WebSocket, a meta-protocol
running on HTTP, but inside the WebSocket framing is
the ordinary TLS-based Tor protocol.
Decoy routing [35] is a technique that puts proxies
in the middle of network paths, rather than at the ends.
For this reason, it is also called end-to-middle proxying. Realizations of decoy routing include Telex [70],
Cirripede [29], and TapDance [69]. Decoy routing asks
friendly ISPs to deploy special routers that reside on
network paths between censored users and uncensored
“decoy” Internet destinations. Circumvention traffic is
“tagged” in a way that is detectable only by the special
routers, and not by the censor. On receiving a tagged
communication, the router shunts it away from its apparent, overt destination and toward a censored, covert
_destination. Domain fronting is similar in spirit to de-_
coy routing: think of domain fronting as decoy routing
at the application layer. In place of a router, domain
fronting has a frontend server; in place of the overt destination is the front domain. Both systems tag flows in
a way that is invisible to the censor: decoy routing uses,
for example, a hash embedded in a client nonce, while
fronting uses the HTTP Host header, encrypted inside
of HTTPS. Fronting has the advantage of not requiring
cooperation by network intermediaries.
Schuhard et al. [54] introduce the idea of a rout_ing adversary against decoy routing, and show that the_
connectivity of the Internet enables censors to force network users onto paths that do not include participating
routers. Simulations by Houmansadr et al. [31] show
that even though such alternate paths exist, they are


-----

many times more costly to the censor, especially when
participating routers are placed strategically.
Collage [11] makes a covert channel out of web sites
that accept user-generated content, like photos. Both
sender and receiver rendezvous through one of these
sites in order to exchange messages. The design of Collage recognizes the need for the proxy sites to be resistant to blocking, which it achieves through the wide
availability of suitable sites.
CloudTransport [10] uses cloud storage services, for
example Amazon S3, as a communication channel by encoding sends and receives as reads and writes to shared
remote files. CloudTransport has much in common with
domain fronting: it hides the true endpoint of a communication using HTTPS, and it sends traffic through a domain with high collateral damage. In CloudTransport,
the hidden destination, which is a storage bucket name
rather than a domain, is hidden in the path component of a URL. For example, in the S3 URL https://s3.
amazonaws.com/bucketname/filename, the censor only
gets to “see” the generic domain part, “s3.amazonaws.
com”. The path component “/bucketname/filename”,
which would reveal the use of CloudTransport, cannot be used for blocking because it is encrypted under
HTTPS.
In a prescient 2012 blog post [8], Bryce Boe described how to gain access to Google HTTPS services
through a single whitelisted Google IP address, by manual editing of a hosts file. He observed that Google App
Engine could serve as a general-purpose proxy, and anticipated the countermeasure of SNI filtering, noting
that sending a false SNI could defeat it.
To our knowledge, the earliest use of domain
fronting for circumvention was by GoAgent [24], a tool
based on Google App Engine and once widely used in
China. Users of GoAgent upload a personal copy of the
proxy code to App Engine, where it runs on a subdomain of appspot.com. In order to reach appspot.com,
GoAgent fronts through a Google IP address, using the
“domainless” model without SNI. GoAgent does not
use an additional general-purpose proxy after fronting;
rather, it fetches URLs directly from the App Engine
servers. Because of that, GoAgent does not support protocols other than HTTP and HTTPS, and the end-toend security of HTTPS is lost, as web pages exist in
plaintext on the App Engine servers before being reencrypted back to the client. According to a May 2013
survey [53], GoAgent was the most-used circumvention
tool in China, with 35% of survey respondents having
used it in the previous month. It ranked higher than
paid (29%) and free VPNs (18%), and far above special

purpose tools like Tor (2.9%) and Psiphon (2.5%).
GoAgent was disrupted in China starting in the beginning of June 2014, when all Google services were
blocked [2, 25]. The block also affected our prototype
systems when used in China with App Engine, though
they continued to work in China over other web services.

## 4 Fronting-capable web services

In this section we survey a variety of web services and
evaluate their suitability for domain fronting. Most of
the services we evaluated support fronting in one form
or another, but they each have their own quirks and
performance characteristics. The survey is not exhaustive but it includes many of the most prominent content
delivery networks. Table 1 is a summary.
Pricing across services varies widely, and depends on
complicated factors such as geographical region, bandwidth tiers, price breaks, and free thresholds. Some services charge per gigabyte or per request, some for time,
and some for other resources. Most services charge between $0.10 and $0.20 per GB; usually bandwidth is
cheaper in North America and Europe than in the rest
of the world.
Recall that even services that support domain
fronting will front only for the domains of their own
customers. Deployment on a new service typically requires becoming a customer, and an outlay of time and
money. Of the services surveyed, we have at some time
deployed on Google App Engine, Amazon CloudFront,
Microsoft Azure, Fastly, and CloudFlare. The others we
have only tested using manually crafted HTTP requests.

**Google App Engine [26] is a web application**
platform. Users can upload a web app needing nothing
more than a Google account. Each application gets a
user-specified subdomain of appspot.com, for which almost any Google domain can serve as a front, including
google.com, gmail.com, googleapis.com, and many others. App Engine can run only web applications serving
short-lived requests, not a general-purpose proxy such
as a Tor bridge. For that reason we use a tiny “reflector”
application that merely forwards incoming requests to
a long-lived proxy running elsewhere. Fronting through
App Engine is attractive in the case where the censor blocks appspot.com but at least one other Google
domain is reachable. App Engine costs $0.12/GB and
$0.05 for each “instance hour” (the number of running
instances of the app is adjusted dynamically to meet
load, and you pay for each instance after the first).


-----

**Fig. 2. Architecture of meek. The client sends an HTTP request to the Tor bridge by way of an intermediate web service such as a**
CDN. The client protects the bridge’s domain name forbidden.example from the censor by fronting it with another name, here al**lowed.example. The intermediate web server decrypts the TLS layer and forwards the request to the bridge according to the Host**
header. The bridge sends data back to the client in the HTTP response. meek-client and meek-server are the interface between Tor
and the pluggable transport; from Tor’s point of view, everything between meek-client and meek-server is an opaque data transport.
The host at allowed.example does not participate in the communication.


**Table 1. Summary of fronting-capable services. The table does**
not include other kinds of services, such as shared web hosting,
that may work for domain fronting. Bandwidth charges usually
vary by geographic region. Many services offer price breaks starting around 10 TB/month. Prices are current as of May 2015 and
are rounded to the nearest cent.

**service** **$/GB** **$/10K reqs.** **$/hour** **$/month**
**App Engine[1]** **0.12** **–** **0.05** **–**
**CloudFront** **0.09–0.25** **0.01–0.02** **–** **–**
**Azure** **0.09–0.14** **–** **–** **–**
**Fastly[2]** **0.12–0.19** **0.01** **–** **–**
**CloudFlare[3]** **–** **–** **–** **200**
**Akamai[4]** **–** **–** **–** **400**
**Level 3[5]** **0.10–0.25** **–** **–** **–**

1 App Engine dynamically scales the number of “instances” of the application code in order to handle
changing load. Every instance after the first costs
$0.05/hour.
2 Fastly has a minimum monthly charge of $50.
3 CloudFlare charges $200/month for its “business”
plan. It has other plans that cost more and less.
4 Akamai does not publish pricing information; the
prices here are from a reseller called Cache Simple,
which quotes $400/month for 1000 GB transfer, and
$0.50/GB for overages.
5 Level 3 does support domain fronting per se, but
paths under the secure.footprint.net domain can
likely serve the same purpose. Level 3 does not publish pricing information; the prices here are from
a reseller called VPS.NET, which quotes $0.10–
0.25/GB.


Applications are free of charge if they stay below certain usage thresholds, for example 1 GB of bandwidth
daily, making possible a distributed, upload-your-ownapp model in the style of GoAgent.
**Amazon CloudFront [3] is the CDN of Ama-**
zon Web Services. A CloudFront “distribution,” as a
CDN configuration is called, associates an automatically generated subdomain of cloudfront.net with an
origin server. The front domain may be any other
cloudfront.net subdomain (all of which support HTTPS
through a wildcard certificate), or any other DNS alias
for them. CloudFront is easy to set up: one must only set
the origin domain and no reflector app is needed. Pricing per GB ranges from $0.085 for the United States
and Europe, up to $0.25 for South America, with price
breaks starting at 10 TB/month. There is an additional charge per 10,000 HTTPS requests, ranging from
$0.0075 in the United States to $0.0160 in South America. CloudFront has a usage tier that is free of charge for
a year, subject to a bandwidth limit of 50 GB/month.
**Microsoft Azure [46] is a cloud computing plat-**
form that features a CDN. Like CloudFront, Azure
assigns automatically generated subdomains of vo.
msecnd.net. any of which can front for any other.
There are other possible front domain names, like ajax.
aspnetcdn.com, that are used as infrastructure by many
web sites, lending them high collateral damage. Unlike CloudFront’s, Azure’s CDN forwards only to Azureaffiliated domains, so as with App Engine, it is necessary
to run a reflector app that forwards requests to some external proxy. Bandwidth costs $0.087–0.138/GB, with
price breaks starting at 10 TB/month.
**Fastly [19] is a CDN. Unlike most CDNs, Fastly**
validates the SNI: if SNI and Host do not match, the


-----

edge server returns an HTTP 400 (“Bad Request”) error. However, if the TLS ClientHello simply omits SNI,
then the Host may be any Fastly domain. Fastly therefore requires the “domainless” fronting style. Fastly’s
pricing model is similar to CloudFront’s. They charge
between $0.12 and $0.19 per GB and $0.0075 and $0.009
per 10,000 requests, depending on the region.
**CloudFlare [12] is a CDN also marketed as pro-**
tection against denial-of-service attacks. Like Fastly,
CloudFlare checks that the SNI matches the Host
header and therefore requires sending requests without
SNI. CloudFlare charges a flat fee per month and does
not meter bandwidth. There is a no-cost plan intended
for small web sites, which is adequate for a personal
domain-fronting installation. The upgraded “business”
plan is $200/month.
**Akamai [1] is a large CDN. Requests may be**
fronted through the special HTTPS domain a248.e.
akamai.net, or other customer-configured DNS aliases,
though it appears that certain special domains get special treatment and do not work as fronts. Akamai has
the potential to provide a lot of cover: in 2010 it carried
15–20% of all web traffic [49]. Akamai does not publish pricing details, but it is reputed to be among the
pricier CDNs. We found a reseller, Cache Simple, that
charges $400 for 1000 GB/month, and $0.50/GB after
that. The special domain a248.e.akamai.net began to
be DNS-poisoned in China in late September 2014 [27]
(possibly because it had been used to mirror blocked
web sites), necessitating an alternative front domain in
that country.
**Level 3 [42] is a tier-1 network operator that has**
a CDN. Unlike other services in this section, Level 3
does not appear to support domain fronting. However,
we mention it because it may be possible to build similar functionality using distinct URL paths under the
domain secure.footprint.net (essentially using the path,
rather than the Host header, as a hidden tag). Level 3
does not publish pricing data. We found a reseller,
VPS.NET, that quotes $34.95 for the first 1000 GB and
$0.10–0.25/GB thereafter. Level 3’s special HTTPS domain secure.footprint.net is also now DNS-poisoned in
China.

There are other potential deployment models apart
from CDNs. For example, there are cheap web hosts
that support both PHP and HTTPS (usually with a
shared certificate). These features are enough to support a reflector app written in PHP, which users can
upload under their own initiative. In this do-it-yourself
model, blocking resistance comes not from a strong front


domain, but from the diffuseness of many proxies, each
carrying only a small amount of traffic. The URLs of
these proxies could be kept secret, or could be carefully disseminated by a proxy-distribution service like
BridgeDB [9]. Psiphon uses this approach when in “unfronted” mode.
Another alternative is deployment with the cooperation of an existing important web site, the blocking
of which would result in high collateral damage. It is a
nice feature of domain fronting that it does not require
cooperation by the intermediate web service, but if you
have cooperation, you can achieve greater efficiency. The
important web site could, for example, reserve a magic
URL path or domain name, and forward matching requests to a proxy running locally. The web site does
two jobs: its ordinary high-value operations that make
it expensive to block, and a side job of handling circumvention traffic. The censor cannot tell which is which
because the difference is confused by HTTPS.

## 5 Deployment on Tor

We implemented domain fronting as a Tor pluggable
transport [5] called meek. meek combines domain
fronting with a simple HTTP-based tunneling proxy.
Domain fronting enables access to the proxy; the proxy
transforms a sequence of HTTP requests into a Tor data
stream.
The components of the system appear in Figure 2.
meek-client acts as an upstream proxy for the client’s
Tor process. It is essentially a web client that knows how
to front HTTPS requests. When meek-client receives
an outgoing chunk of data from a client Tor process,
it bundles the data into a POST request and fronts the
request through the web service to a Tor bridge. The Tor
bridge runs a server process, meek-server, that decodes
incoming HTTP requests and feeds their data payload
into the Tor network.
The server-to-client stream is returned in the bodies of HTTP responses. After receiving a client request,
meek-server checks for any pending data the bridge
has to send back to the client, and sends it back in
the HTTP response. When meek-client receives the response, it writes the body back into the client Tor.
The body of each HTTP request and response carries a small chunk of an underlying TCP stream (up
to 64 KB). The chunks must be reassembled, in order,
without gaps or duplicates, even in the face of transient failures of the intermediate web service. meek uses


-----

```
POST / HTTP/1.1
Host: forbidden.example
X-Session-Id: cbIzfhx1HnR
Content-Length: 517
\x16\x03\x01\x02\x00\x01\x00\x01\xfc\x03\x03\x9b\xa9...
  HTTP/1.1 200 OK
  Content-Length: 739
  \x16\x03\x03\x00\x3e\x02\x00\x00\x3a\x03\x03\x53\x75...
POST / HTTP/1.1
Host: forbidden.example
X-Session-Id: cbIzfhx1HnR
Content-Length: 0
  HTTP/1.1 200 OK
  Content-Length: 75
  \x14\x03\x03\x00\x01\x01\x16\x03\x03\x00\x40\x06\x84...

```
**Fig. 3. Requests and responses in the meek HTTP proto-**
col. The session ID is randomly generated by the client. Request/response bodies contain successive chunks of a Tor TLS
stream (\x16\x03\x01 is the beginning of a TLSv1.0 ClientHello
message). The second POST is an empty polling request. The
messages shown here are encrypted inside HTTPS until after
they have been fronted, so the censor cannot use the Host and
X-Session-Id headers for classification.

a simple approach: requests and responses are strictly
serialized. The client does not send a second chunk of
data (i.e., make another request) until it has received
the response to its first. The reconstructed stream is
simply the concatenation of bodies in the order they
arrive. This technique is simple and correct, but less efficient because it needs a full round-trip between every
send. See Sections 6 and 7 for alternative approaches
that increase efficiency.
meek-server must be able to handle many simultaneous clients. It maintains multiple connections to a
local Tor process, one for each active client. The server
maps client requests to Tor connections by “session ID,”
a token randomly generated by the client at startup.
The session ID plays the same role in the meek protocol
that the (source IP, source port, dest IP, dest port) tuple
plays in TCP. The client sends its session ID in a special
X-Session-Id HTTP header. meek-server, when it sees a
session ID for the first time, opens a new connection
to the local Tor process and adds a mapping from ID
to connection. Later requests with the same session ID
reuse the same Tor connection. Sessions are closed after
a period of inactivity. Figure 3 shows a sample of the
protocol.
HTTP is fundamentally a request-based protocol.
There is no way for the server to “push” data to the
client without having first received a request. In order


Azure

Google

CloudFront


14 s 48 s

21 s 60 s

31 s 101 s

0 30 60 90
Time (seconds)

|Col1|21 s|Col3|6|0 s|Col6|Col7|
|---|---|---|---|---|---|---|
||3|1 s||||10|


**Fig. 4. Time to download an 11 MB file through Tor, with and**
without meek. Text labels indicate the mean of 10 measurements.
Bulk-download times increase by about a factor of 3 when meek
is activated. The middle and exit nodes are constant across all
measurements; only the entry varies according to the service. The
without-meek circuits use an entry node located at the same IP
address as the corresponding with-meek circuits.

to enable the server to send back data, meek-client sends
occasional empty polling requests even when it has no
data to send. The polling requests simply give the server
an opportunity to send a response. The polling interval
starts at 100 ms and grows exponentially up to a maximum of 5 s.
The HTTP-based tunneling protocol adds overhead. Each chunk of data gets an HTTP header, then
the HTTP request is wrapped in TLS. The HTTP
header adds about 160 bytes [59], and TLS adds another 50 bytes or so (the exact amount depends on
the ciphersuite chosen by the intermediate web service).
The worst-case overhead when transporting a single encrypted Tor cell of about 540 bytes is about 40%, and
less when more than one cell is sent at once. We can
estimate how much overhead occurs in practice by examining CDN usage reports. In April 2015, the Amazon CloudFront backend for meek received 3,231 GB
in 389 M requests [21], averaging about 8900 bytes
per request. If the overhead per request is 210 bytes,
then the average overhead is 210/(8900 − 210) ≈ 2.4%.
meek-client reuses the same TLS connection for many
requests, so the TLS handshake’s overhead is amortized.
Polling requests also use bandwidth, but they are sent
only when the connection is idle, so they do not affect
upload or download speed.
Figure 4 measures the effect of meek’s overhead on
download speed. It shows the time taken to download
a 11,536,384-byte file (http://speedtest.wdc01.softlayer.
com/downloads/test10.zip) with and without meek,
over the three web services on which we have deployed.
We downloaded the file 10 times in each configuration.


-----

**App Engine** **CloudFront** **Azure (est.)**
**GB** **cost** **GB** **cost** **GB** **cost**
**early 2014** **71** **$8** **67** **$8** **47** **$5**
**Oct 2014** **289** **$41** **479** **$130** **296** **$31**
**Nov 2014** **1,375** **$225** **1,269** **$363** **499** **$53**
**Dec 2014** **2,132** **$327** **1,579** **$417** **511** **$64**
**Jan 2015** **2,944** **$464** **2,449** **$669** **637** **$68**
**Feb 2015** **4,114** **$651** **2,369** **$605** **614** **$65**
**Mar 2015** **5,316** **$690** **3,385** **$816** **736** **$78**
**Apr 2015** **6,304** **$886** **3,231** **$785** **1,982** **$210**
**total** **22,545 $3,292** **14,828 $3,793** **5,328** **$565**

**Fig. 5. Concurrent users of the meek pluggable transport with**
month-by-month transfer and cost. The Azure columns are estimates of what we would pay if we did not have a special research
grant. User counts come from the Tor Metrics Portal [43, 60].

The time to download the file increases by about a factor of 3 when meek is in use. We attribute this increase
to the added latency of an indirect path through the
CDN, and the latency-bound nature of meek’s naive serialization.
meek’s primary deployment vehicle is Tor
Browser [51], a derivative of Firefox that is preconfigured to use a built-in Tor client. Tor Browser features
an easy interface for enabling meek and other pluggable
transports. Deployment began in earnest in October
2014 with the release of Tor Browser 4.0 [50], the first
release to include meek as an easy selectable option.
It runs over Google App Engine, Amazon CloudFront,
and Microsoft Azure. Figure 5 shows the daily average
number of concurrent users. (A value of 1,000, for example, means that there were on average 1,000 users of
the system at any time during the day.) Also in Figure 5 is a table of monthly costs broken down by web
service. Our Azure service is currently running on a free
research grant, which does not provide us with billing
information. We estimate what Azure’s cost would
be by measuring the bandwidth used at the backing
Tor bridge, and assuming bandwidth costs that match
the geographic traffic mix we observe for CloudFront:
roughly 62% from North America and Europe, and 38%
from other regions.


### 5.1 Camouflage for the TLS layer

Without additional care, meek would be vulnerable to
blocking by its TLS fingerprint. TLS, on which HTTPS
is based, has a handshake that is largely plaintext [13,
§7.4] and leaves plenty of room for variation between
implementations. These differences in implementation
make it possible to fingerprint TLS clients [44]. Tor itself was blocked by China in 2011 because of the distinctive ciphersuites it used at the time [57]. Figure 12a in
Appendix A shows how meek-client’s fingerprint would
appear natively; it would be easy to block because not
much other software shares the same fingerprint. Figures 12b and 12c show the fingerprints of two web
browsers, which are more difficult to block because they
also appear in much non-circumvention traffic.
In order to disguise its TLS fingerprint, meek-client
proxies all its HTTPS requests through a real web
browser. It looks like a browser, because it is a browser.
We wrote extensions for Firefox and Chrome that enable
them to make HTTPS requests on another program’s
behalf. The browser running the extension is completely
separate from the Tor Browser the user interacts with.
It runs in the background in a separate process, does
not display a user interface, and shares no state with
the user’s browser. The extra cost of this arrangement
is negligible in terms of latency, because communication
with the headless browser occurs over a fast localhost
connection, and in terms of CPU and RAM it is the
same as running two browsers at once.
The client’s Tor process starts both meek-client and
the headless browser, then configures meek-client to
proxy its requests through the browser. The headless
browser is the only component that actually touches
the network. It should be emphasized that the headless browser only makes domain-fronted requests to the
front domain; the URLs it requests have no relation to
the pages the user browses.

## 6 Deployment on Lantern

Lantern [40] is a free circumvention tool for casual web
browsing. It does not employ onion routing and focuses more on performance and availability than on
anonymity. Lantern encompasses a network of shared
HTTPS proxy servers, and client software that allows
censored users to find and use those proxy servers with
their existing web browsers. The Lantern client also
allows uncensored users to host proxy servers (“peer

-----

**Fig. 6. MB/s served by domain-fronted Lantern proxies (both**
Lantern-hosted and peers).

hosted” servers) for use by others. Lantern aims to
provide a secure mechanism for distributing knowledge about both Lantern-hosted and peer-hosted proxy
servers using a trust network–based distribution mechanism such as Kaleidoscope [56]. In the meantime,
Lantern also randomly assigns users to Lantern-hosted
proxy servers.
Lantern has a centralized infrastructure for authenticating users and assigning them proxies. Its threat
model assumes that the centralized infrastructure may
be blocked by censors. Therefore, users must have a priori access to an unblocked proxy (a “fallback”) which
they use to bootstrap into the rest of the network.
Lantern originally distributed the IP addresses of
fallbacks by embedding them in customized software installers that we sent to users via email autoresponder.
This method prevented users from directly downloading
Lantern from our website and would have made it easy
for censors to discover proxies simply by signing up for
Lantern (though in practice we never saw this happen).

### 6.1 Implementation

We rolled out domain fronting in July 2014, allowing
users to download Lantern directly for the first time.
The directly downloaded clients proxied all their traffic
via domain fronting. After initial testing with Fastly, we
changed to a different CDN, which has proven attractive
because it has many unblocked front domains, it does
not charge for bandwidth, and its API enables us to
easily register and unregister proxies.
Figure 6 shows user bandwidth since deployment.
After experiencing steady growth, in October 2014 we
started randomly assigning direct HTTPS proxies to
users who had direct-downloaded Lantern. This diverted some traffic from domain fronted servers to more
efficient direct servers. In December 2014 and January 2015, there was a dramatic surge in domain-fronted
traffic, which jumped from 1 MB/s to 100 MB/s within


those two months. Activity has remained at around the
100 MB/s level since then.
Lantern’s domain fronting support is provided by
an application called flashlight [38], which uses library
layers called enproxy [37] and fronted [39]. enproxy provides an abstract network connection interface that encodes reads and writes as a sequence of HTTP requests
via a stateful enproxy proxy. enproxy allows flashlight
to proxy any streaming-oriented traffic like TCP. Unlike
Tor’s implementation of meek, enproxy supports fullduplex transfer, which is handy for bidirectional protocols like XMPP, which Lantern uses for P2P signaling. fronted uses domain fronting to transmit enproxy’s
HTTP requests in a blocking-resistant manner. In practice, we configure fronted with several hundred host domains that are dialed via IP address (no DNS lookup).
Domain-fronted Lantern requests go to domain
names, such as fallbacks.getiantem.org, that represent
pools of servers. The CDN distributes requests to the
servers in round-robin fashion. The domain-fronting
protocol is stateful, so subsequent HTTP requests for
the same connection are routed to the original responding proxy using its specific hostname (sticky routing),
which the client obtains from a custom HTTP header.
The proxy hostname serves the same request-linking
purpose as the session ID does in meek.

### 6.2 Mitigations for increased latency

The encoding of a stream as a sequence of HTTP requests introduces additional latency beyond that of
TCP. In the case of flashlight with our chosen CDN,
the additional latency has several causes. We describe
the causes and appropriate mitigations.
Domain fronting requires the establishment of additional TCP connections. The client, the CDN, and
the proxy between themselves introduce three additional TCP connections between the client and the destination. To reduce latency, the CDN pools and reuses
connections to the Lantern proxy. Unfortunately, the
Lantern client cannot do the same for its connections
to the CDN because the CDN seems to time out idle
connections fairly aggressively. We mitigate this by aggressively pre-connecting to the CDN when we detect
activity [36].
Though enproxy is mostly full duplex, reads cannot begin until the first request and its response with
the sticky-routing header have been processed. This is
a basic limitation.


-----

direct

non-fronted proxy

fronted proxy


5 s

6 s

9 s

0.0 2.5 5.0 7.5 10.0 12.5
Time (seconds)

|Col1|Col2|Col3|Col4|6|s|Col7|Col8|Col9|Col10|
|---|---|---|---|---|---|---|---|---|---|
||||||||9 s|||


**Fig. 7. Time to download an 11 MB file direct, through a one-**
hop proxy, and through Lantern with domain fronting. The enproxy transport and extra CDN hop increase download times by
about a factor of 2. Text labels indicate the mean of 10 measurements.

enproxy does not pipeline HTTP requests. Even
though reads and writes are full duplex, a write cannot proceed until the flush of previous writes has been
acknowledged with an HTTP response—a full roundtrip is necessary between each flush. In the future,
HTTP/2’s request pipelining will potentially improve
on latency.
enproxy has no way of knowing when the client is
done writing. If the data were streaming directly from
the client through to the proxy, this would not be a
problem, but CDNs buffer uploads: small uploads aren’t
actually forwarded to the proxy until the HTTP request
is finished. enproxy assumes that a writer is finished
if it detects inactivity for more than 35 ms, at which
point it flushes the write by finishing the HTTP request.
This introduces at least 35 ms of additional latency, and
potentially more if the guess is wrong and the write is
not actually finished, since we now have to wait for a
full round trip before the next write can proceed. This
latency is particularly noticeable when proxying TLS
traffic, as the TLS handshake consists of several small
messages in both directions.
This last source of latency can be eliminated if enproxy can know for sure when a writer is finished. This
could be achieved by letting enproxy handle the HTTP
protocol specifically. Doing so would allow enproxy to
know when the user agent is finished sending an HTTP
request and when the destination is finished responding.
However, doing the same for HTTPS would require a
local man-in-the-middle attack on the TLS connection
in order to expose the flow of requests. Furthermore,
this approach would work only for HTTP clients. Other
traffic, like XMPP, would require additional support for
those protocols.
Figure 7 compares download speeds with and
without a domain fronting proxy. It is based on
10 downloads of the same 11 MB file used in the


Tor bandwidth test in the previous section, however
located on a different server close to the Lantern
proxy servers: http://speedtest.ams01.softlayer.com/
downloads/test10.zip. The fronted proxy causes download times to approximately double. Because of
Lantern’s round-robin rotation of front domains, the
performance of the fronted proxy may vary over time
according to the route to the CDN.

### 6.3 Direct domain fronting

The Lantern network includes a geolocation server. This
server is directly registered on the CDN and the Lantern
client domain-fronts to it without using any proxies, reducing latency and saving proxy resources. This sort of
direct domain fronting technique could in theory be implemented for any web site simply by registering it under
a custom domain such as facebook.direct.getiantem.org.
It could even be accomplished for HTTPS, but would
require the client software to man-in-the-middle local
HTTPS connections between browser and proxy, exposing the plaintext not only to the Lantern client but also
to the CDN. In practice, web sites that use the CDN
already expose their plaintext to the CDN, so this may
be an acceptable solution.

## 7 Deployment on Psiphon

The Psiphon circumvention system [52] is a centrally
managed, geographically diverse network of thousands
of proxy servers. It has a performance-oriented, one-hop
architecture. Much of its infrastructure is hosted with
cloud providers. As of January 2015, Psiphon has over
two million daily unique users. Psiphon client software
runs on popular platforms, including Windows and Android. The system is designed to tunnel a broad range of

300,000

200,000

100,000

0

Dec 15 Jan 01 Jan 15 Feb 01
2014 2015 2015 2015

**Fig. 8. Daily unique users of meek with Psiphon. Clients send a**
coarse-grained “last connected” timestamp when connecting. A
unique user is counted whenever a user connects with a timestamp before the day under consideration.


-----

host traffic: web browsing, video streaming, and mobile
app data transfer. Client software is designed for ease of
use; users are not asked to perform any configuration.
Psiphon has faced threats including blocking by
DPI—both blacklisting and whitelisting—and blocking
by address. For example, in 2013, Psiphon circumvented
HTTP-whitelisting DPI by sending an “HTTP prefix”
(the first few bytes of an HTTP request) before the start
of its regular upstream flow [6].
Psiphon strives to distribute its server addresses in
such a way that most clients discover enough servers
to have several options in the case of a server being blocked, while making it difficult to enumerate all
servers. In February 2014, Psiphon was specifically targeted for address-based blocking, and this blocking was
aggressive enough to have a major impact on our user
base, though not all users were blocked. As part of
our response we integrated and deployed meek-based
domain fronting, largely based on Tor’s implementation, with some modifications. It was fully deployed in
June 2014. Figure 8 shows the number of unique daily
users of fronted meek with Psiphon.
In addition, Psiphon also employs meek in what
we call “unfronted” mode. Unfronted meek omits the
TLS layer and the protocol on the wire is HTTP. As
fully compliant HTTP, unfronted meek supersedes the
“HTTP prefix” defense against HTTP whitelisting. Unfronted meek is not routed through CDNs, and as such is
only a defense against DPI whitelisting and not against
proxy address enumeration. We envision a potential future fronted HTTP protocol with both properties, which
requires cooperating with CDNs to route our HTTP requests based on, for example, some obfuscated HTTP
header element.

### 7.1 Implementation

Psiphon’s core protocol is SSH. SSH provides an encryption layer for communication between Psiphon clients
and servers; the primary purpose of this encryption is
to frustrate DPI. On top of SSH, we add an obfuscatedopenssh [41] layer that transforms the SSH handshake
into a random stream, and add random padding to the
handshake. The payload within the meek transport appears to be random data and lacks a trivial packet size
signature in its initial requests and responses. Psiphon
clients authenticate servers using SSH public keys obtained out of band, a process that is bootstrapped with
server keys embedded in the client binaries.


Psiphon uses a modified version of the meek protocol described in Section 5. The session ID header contains extra information: a protocol version number and
the destination Psiphon server address. As this cookie
will be visible to the censor in unfronted mode, its
value is encrypted in a NaCl crypto_box [7] using the
public key of the destination meek-server; then obfuscated; then formatted as an innocuous-seeming cookie
with a randomly selected key. meek-server uses the protocol version number to determine if the connecting
meek-client supports Psiphon-specific protocol enhancements. The destination address is the SSH server to
which meek-server should forward traffic.
In Psiphon, meek-client transmits its chosen session ID on its first HTTP request, after which
meek-server assigns a distinct ID to be used on subsequent requests. This change allows meek-server to distinguish new and existing sessions when a client sends a
request after a long delay (such as after an Android device awakes from sleeping), when meek-server may have
already expired and discarded its session.
We ported meek-client, originally written in Go, to
Java for Android. On Android, we make HTTP and
HTTPS requests using the Apache HttpClient component, in order to have a TLS fingerprint like those of
other Android apps making web service requests.
The Psiphon meek-server inspects CDN-injected
headers, like X-Forwarded-For, to determine the client’s
IP address. The address is mapped to a geographic region that is used in recording usage statistics.

### 7.2 Server selection

When a user starts a Psiphon client, the client initiates connections to up to ten different servers simultaneously, keeping the first to be fully established. Candidate servers are chosen at random from cached lists
of known servers and a mix of different protocols, both
fronted and non-fronted, are used. The purpose of the
simultaneous connections is to minimize user wait time
in case certain protocols are blocked, certain servers are
blocked by address, or certain servers are at capacity
and rejecting new connections. This process also tends
to pick the closest data center, and the one with lowest
cost, as it tends to pick lower-latency direct connections
over domain-fronted connections.
We made two modifications to server selection in
order to accommodate fronting. First, we changed the
notion of an “established connection” from TCP connection completion to full SSH handshake completion.


-----

without meek

meek (streaming)

meek (no streaming)


5 s

11 s

21 s

0 10 20
Time (seconds)

|Col1|1|1 s|Col4|Col5|
|---|---|---|---|---|
||||2|1 s|


**Fig. 9. Time to download an 11 MB file through Psiphon over**
three transports: one-hop obfuscated-openssh proxy without
meek; meek with streaming downloads; and meek without
streaming downloads. Text labels indicate the mean of the 50
fastest measurements out of 100. Bulk-download times increase
by about a factor of 3 when meek is activated. With fixed-size
HTTP bodies, meek costs about a factor of 4 in download time.
With the optimization of unlimited-size HTTP bodies, the download time decreases to about a factor of 2.

This ensures that both hops are measured in the fronted
case. Second, we adjusted our protocol selection schedule to ensure that, while we generally favor the fastest
connection, we do not expose the system to an attack
that would force us to use a degraded protocol. For example, a censor could use a DPI attack that allows all
connections to establish, but then terminate or severely
throttle non-whitelisted protocols after some short time
period. If the client detects such degraded conditions,
it begins to favor fronted and unfronted protocols over
the faster obfuscated SSH direct connections.

### 7.3 Performance

We identified a need to improve the video streaming and
download performance of meek-tunneled traffic. In addressing this, we considered the cost per HTTP request
of some candidate CDNs, a lack of support for HTTP
pipelining in our components, and a concern about the
DPI signature of upstream-only or downstream-only
HTTP connections. As a compromise between these
considerations, we made a tweak to the meek protocol: instead of sending at most 64 KB in each HTTP
response, responses stream as much as possible, as long
as there is data to send and for up to 200 ms.
This tweak yielded a significant performance improvement, with download speeds increasing by up to
4–5×, and 1080p video playback becoming smooth. Under heavy downstream conditions, we observe response
bodies up to 1 MB, 300 KB on average, although the exact traffic signature is highly dependent on the tunneled
application. We tuned the timeout parameter through


subjective usability testing focused on latency while web
browsing and simultaneously downloading large files.
Figure 9 compares the time taken to download a
file both with and without meek, and with and without
the streaming download optimization. The target is the
same speedtest.wdc01 URL used in the Tor performance
tests in Section 5. The performance effect of meek is
about a factor-4 increase in download time; streaming
downloads cut the increase in half.

## 8 Traffic analysis

In developing domain fronting circumvention systems,
we hope to deprive the censor of easy distinguishers
and force the use of more expensive, less reliable classification tests—generally, to increase the cost of censorship. We believe that domain fronting, implemented
with care, meets the primary challenges of proxy-based
circumvention. It defeats IP- and DNS-based blocking
because the IP- and DNS-layer information seen by the
censor are not those of the proxy; content-based blocking because content is encrypted under HTTPS; and
active probing because though a censor may be able to
discover that a web service is used for circumvention,
it cannot block the service without incurring significant
collateral damage.
Our experiences with deploying circumvention systems has led us to conclude that other potential means
of censorship—e.g., identifying circumventing content
by analyzing packet length distributions—do not currently have relevance when considering the practices of
today’s censors. We speculate that censors find such
tests unattractive because they require storing significant state and are susceptible to misclassification. More
broadly, we are not aware of any nation-level censorship
event that made use of such traffic features.
Nevertheless, we expect censors to adapt to a changing environment and to begin deploying more sophisticated (but also more expensive and less reliable) tests.
The issue of traffic analysis is a general one [68], and
mostly separable from domain fronting itself. That is,
domain fronting does not preclude various traffic shaping techniques and algorithms, which can be developed
independently and plugged in when the censors of the
world make them necessary. This section contains a
sketch of domain fronting’s resistance to certain traffic analysis features, though a case study of meek with
Tor and a trace of non-circumvention traffic. While we
identify some features that may give a censor leverage


-----

1.00

0.75

0.50

0.25

0.00

|Col1|Col2|Col3|Col4|Col5|Col6|Col7|Col8|Col9|Col10|Col11|Col12|
|---|---|---|---|---|---|---|---|---|---|---|---|
|||LBL Goo|gle H|TTP|S|||||||
|||||||||||||
|||||||||||||
|||||||||||||
||||me|ek o|n|Ap|p Engi|ne||||
|||||||||||||
|||||||||||||
|||||||||||||


5 10 24 60 120180 300 600900 1800 3600
Duration (seconds)


**LBL Google HTTPS** **meek on App Engine**
**0 bytes** **37.6%** **1418 bytes** **40.5%**
**1430 bytes** **9.1%** **0 bytes** **37.7%**
**1418 bytes** **8.5%** **1460 bytes** **7.2%**
**41 bytes** **6.1%** **396 bytes** **2.0%**
**1416 bytes** **3.1%** **196 bytes** **1.8%**
**1460 bytes** **2.9%** **1024 bytes** **1.5%**

**Fig. 10. Comparison of TCP payload length distributions in or-**
dinary HTTPS connections to Google services from the LBL
traffic trace, and meek running on App Engine, fronted through
www.google.com.


**Fig. 11. CDF of connection duration. The x-axis is logarithmic.**

the two are not grossly different. In both cases, about
38% of packets are empty (mostly ACKs), with many
packets near the usual TCP Maximum Segment Size of
1460 bytes. Conspicuous in the meek trace are a small
peaks at a few specific lengths, and a lack of short payloads of around 50 bytes. Both of characteristics are
probably reflections of the fixed cell size of the underlying Tor stream.


in distinguishing circumvention traffic, we believe that
the systems we have deployed are sufficiently resistant
to the censors of today, and do not block the way to
future enhancements to traffic analysis resistance.
As domain fronting is based on HTTPS, we evaluate
distinguishability from “ordinary” HTTPS traffic. We
compare two traffic traces. The first is HTTPS traffic
from Lawrence Berkeley National Laboratory (LBL), a
large (≈ 4K users) research lab, comprising data to and
from TCP port 443 on any Google server. Its size is
313 MB (packet headers only, not payloads) and it lasts
10 minutes. The IP addresses in this first trace were
masked, replaced by a counter. The second trace is of
meek in Tor Browser, browsing the home pages of the
top 500 Alexa web sites over Google and App Engine.
It is 687 MB in size and covers 4.5 hours.

### 8.1 Packet length distribution


A censor could attempt to block an encrypted tunnel
by its distribution of packet lengths, if it is distinctive
enough. Figure 10 compares the packet length distributions of the sample traces. Keeping in mind that the
LBL trace represents many users, operating systems,
and web browsers, and the meek trace only one of each,


### 8.2 Connection lifetime

The total duration of TCP connections is another potential distinguisher. Figure 11 shows the cumulative
probability of connection durations in the two traces.
The LBL trace has interesting concentrations on certain round numbers: 10/60/120/180/240 seconds. We
hypothesize that they are caused by keepalive timeouts
in web browsers and servers and periodic polling by web
apps. The small rise at 600 seconds is an artifact caused
by the 10-minute duration of the trace. We do not know
how much longer than 10 minutes those connections
lasted, but they are only 8% of observed connections.
The meek trace shows a propensity for longer connections. In 4.5 hours, there were only 10 connections,
three of them lasting for an hour. The long connections are caused by the client browser extension’s aggressive use of long-lived HTTP keepalive connections,
and by its being constantly busy, giving every opportunity for connection reuse. 60% of meek’s connections
lasted five minutes or longer, while only 13% of ordinary
traffic’s did. meek had essentially no connections lasting
less than 24 seconds, but such short connections were
over 42% of the LBL trace. 30% (3 out of 10) of meek’s
connections lasted almost exactly one hour, evidently
reflecting a built-in keepalive limit in either the client
browser extension or in App Engine.
In light of these measurements, the censor may decide simply to terminate long-lived HTTPS connections.


-----

According to our traffic trace, doing so will not disrupt
more than 8% of ordinary connections (although such
long connections may be valuable large transfers with
higher collateral damage). The censor can lower the timing threshold, at the cost of more false positives. In order
to be effective, then censor must cut off the client completely; otherwise the client may start a new connection
with the same session ID and begin where it left off.

We do not know of any obvious traffic characteristics that reliably distinguish domain fronting from
other HTTPS traffic. Long-lived connections and packet
lengths are potential targets for a more concerted attack. We are fundamentally trying to solve a problem
of steganography, to make circumvention traffic fit some
model of “normal” traffic. However, this can be regarded
as an advantage. What is a challenge for the evaluator
is also a challenge for the censor, simply because it is
difficult to characterize just what normal traffic is, especially behind a CDN that may host variety of services
such as software updates, video streaming, and ordinary web pages. Circumvention traffic need not be perfectly indistinguishable, only indistinguishable enough
that that blocking it causes more and costlier false positives than the censor can accept.

## 9 Discussion

Domain fronting derives its strength from the collateral damage that results from blocking the front domain. It should not—nor should any other circumvention technique—be thought of as unblockable; rather,
one should think of what it costs the censor to block
it. What is unblockable by one censor may be blocked
by another that has different resources and incentives.
Blocking resistance depends on the strength of the front
domain and on the censor’s cost calculus, which has
both economic and social components.
We can at least roughly quantify the cost of blocking any domain fronting system in general. It is the
minimum cost of: blocking a domain; deploying traffic
analysis to distinguish circumvention from other traffic;
or conducting some attack outside our threat model,
for example physical surveillance of Internet users. A
censor could also, for example, block HTTPS entirely,
but that is likely to be even more damaging than targeted blocking of a domain. The cost of blocking a
domain—and the benefit of blocking circumvention—
will vary by censor. For example, China can afford to
block twitter.com and facebook.com partly because it


has domestic replacements for those services, but not
all censors have the same resources. In June 2014, the
Great Firewall of China took the unprecedented step
of blocking all Google services [2, 25], including all potential fronts for App Engine. It is not clear whether
the blocking targeted domain fronting systems like GoAgent; our own systems were only prototypes at that
point. Since then, domain fronting to App Engine has
been effectively stamped out in China, though it continues to work over other web services.
A censor could directly confront the operators of an
intermediate web service and ask them to disable domain fronting (or simply get rid of customers like us who
facilitate circumvention). The censor could threaten to
block the service entirely, costing it business. Whether
such an attack succeeds again depends on specific costs
and motivations. A powerful censor may be able to carry
out its threat, but others will harm themselves more by
blocking a valuable service than the circumvention traffic is worth.
Reliance on paid web services creates the potential for a “financial denial of service” attack against
domain fronting systems, in which the censor uses the
service excessively in an attempt to drive up the operators’ costs. In March 2015, the anticensorship group
GreatFire, which had used various cloud services for
censorship circumvention in China, was the target of
a distributed denial of service attack against their hosting on Amazon Web Services [55]. The attack lasted
for days and incurred tens of thousands of dollars in
bandwidth charges. The attack against Amazon was followed shortly by one against GitHub, the largest in
the site’s history [48]. The second attack specifically
targeted GreatFire’s accounts there. The available evidence indicates that both attacks were coordinated from
within China, using an offensive network system dubbed
the “Great Cannon” [45]. Such an attack could be mitigated by active defenses that shut down a service when
it is being used excessively, though this only protects
against ruinous costs and will not defeat a long-term
attack. It is noteworthy that a world-class censor’s first
reaction was a disruptive, unsubtle denial of service
attack—though we cannot say for sure that the censor did not have something better up its sleeve. GreatFire speculated that the attacks were precipitated by
the publication of an article in the Wall Street Jour_nal [16] that described in detail domain fronting and_
other “collateral freedom” techniques. The interview associated with the article also caused CloudFlare to begin
matching SNI and Host header, in an apparent attempt
to thwart domain fronting.


-----

Fronting shares a potential weakness with decoy
routing, which is that the network paths to the overt and
covert destinations diverge. The difference in paths may
create side channels—different latencies for instance—
that distinguish domain-fronted traffic from the traffic
that really arrives at its apparent destination. For example, a CDN can be expected to have responses to
some fraction of requests already in cache, and respond
to those requests with low latency, while domain-fronted
requests always go all the way to the destination with
higher latency. Schuhard et al. [54, §5] applied latency
measurement to decoy routing. The authors of TapDance [69, §5.1] observe that such an attack is difficult
to carry out in practice, because it requires knowledge
of the performance characteristics of many diverse resources behind the proxy, some of which are not accessible to the censor (login-protected web pages, for example). Domain fronting favors the circumventor even
more, because of the variety of resources behind a CDN.
The intermediate web service has a privileged network position from which it may monitor domainfronted traffic. Even though the censor does not know
which client IP addresses are engaging in circumvention, the CDN knows. The risk is especially acute when
client browses a web site of the same entity that controls the intermediate web server, for example browsing YouTube while fronting through www.google.com.
When this happens, the web service gets to see both
entry and exit traffic, and is in a better position to attempt to correlate flows by timing and volume, even
when the underlying channel is an encrypted protocol
like Tor. This phenomenon seems hard to counter, because the front domain needs to be a popular one in
order to have high collateral damage, but popular domains are also the ones that users tend to want to visit.
It is in theory possible to dynamically switch between
multiple fronts, so as to avoid the situation where the
destination and front are under the same control, at the
cost of leaking information about where the user is not
going at a given moment.
A censor that can man-in-the-middle HTTPS connections can detect domain fronting merely by removing
encryption and inspecting the Host header. Unless the
censor controls a certificate authority, this attack falls to
ordinary HTTPS certificate validation. Against a censor
that controls a trusted certificate authority, certificate
pinning is an effective defense. If the underlying transport is an authenticated and encrypted one like Tor,
then the destination and contents of a user’s connection will remain secret, even if the user is outed as a
circumventor.


## 10 Summary

We have presented domain fronting, an applicationlayer censorship circumvention technique that uses different domain names at different layers of communication in order to hide the true destination of a message.
Domain fronting resists the main challenges offered by
the censors of today: content blocking, address blocking, and active probing. We have implemented domain
fronting in three popular circumvention systems: Tor,
Lantern, and Psiphon, and reported on the experience
of deployment. We begin an investigation into the more
difficult, less reliable means of traffic analysis that we
believe will be necessary to block domain fronting.

## Code and acknowledgments

[The meek pluggable transport has a home page at https:](https://trac.torproject.org/projects/tor/wiki/doc/meek)
[//trac.torproject.org/projects/tor/wiki/doc/meek and](https://trac.torproject.org/projects/tor/wiki/doc/meek)
[source code at https://gitweb.torproject.org/pluggable-](https://gitweb.torproject.org/pluggable-transports/meek.git)
[transports/meek.git. The source code of Lantern’s](https://gitweb.torproject.org/pluggable-transports/meek.git)
[flashlight proxy is at https://github.com/getlantern/](https://github.com/getlantern/flashlight)
[flashlight; other components are in sibling reposito-](https://github.com/getlantern/flashlight)
[ries. Psiphon’s source code is at https://bitbucket.org/](https://bitbucket.org/psiphon/psiphon-circumvention-system)
[psiphon/psiphon-circumvention-system.](https://bitbucket.org/psiphon/psiphon-circumvention-system)
We would like to thank Yawning Angel, George Kadianakis, Georg Koppen, Lunar, and the members of
the tor-dev, tor-qa, and traffic-obf mailing lists who
responded to our design ideas, reviewed source code,
and tested our prototypes. Arlo Breault wrote the
flashproxy-reg-appspot program mentioned in Section 3,
an early application of domain fronting. Leif Ryge and
Jacob Appelbaum tipped us off that domain fronting
was possible. Sadia Afroz, Michael Tschantz, and Doug
Tygar were sources of inspiring conversation. Johanna
Amann provided us with an estimate of the fraction of
SNI-bearing TLS handshakes.
This work was supported in part by the National
Science Foundation under grant 1223717. The opinions,
findings, and conclusions expressed herein are those of
the authors and do not necessarily reflect the views of
the sponsors.


-----

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_conference on Computer and Communications Security_
_[(CCS), Oct. 2012. http://www-users.cs.umn.edu/~hopper/](http://www-users.cs.umn.edu/~hopper/decoy-ccs12.pdf)_
[decoy-ccs12.pdf.](http://www-users.cs.umn.edu/~hopper/decoy-ccs12.pdf)

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[918.pdf.](http://kscope.news.cs.nyu.edu/pub/TR-2008-918.pdf)

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[list, Dec. 2011. https://bugs.torproject.org/4744.](https://bugs.torproject.org/4744)

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[2013. https://bugs.torproject.org/8860.](https://bugs.torproject.org/8860)

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[html?graph=userstats-bridge-transport&end=2015-05-15&](https://metrics.torproject.org/userstats-bridge-transport.html?graph=userstats-bridge-transport&end=2015-05-15&transport=meek)
[transport=meek.](https://metrics.torproject.org/userstats-bridge-transport.html?graph=userstats-bridge-transport&end=2015-05-15&transport=meek)

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[html?graph=userstats-bridge-transport&end=2015-05-15&](https://metrics.torproject.org/userstats-bridge-transport.html?graph=userstats-bridge-transport&end=2015-05-15&transport=obfs3)
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[pdf.](https://netfiles.uiuc.edu/qwang26/www/publications/censorspoofer.pdf)

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[Wustrow.pdf.](https://www.usenix.org/events/sec/tech/full_papers/Wustrow.pdf)


-----

## A Sample TLS fingerprints


**(a) Go 1.4.2’s crypto/tls library**
```
Ciphersuites (13):
 TLS_ECDHE_RSA_WITH_AES_128_GCM_SHA256
 TLS_ECDHE_ECDSA_WITH_AES_128_GCM_SHA256
 TLS_ECDHE_RSA_WITH_RC4_128_SHA
 TLS_ECDHE_ECDSA_WITH_RC4_128_SHA
 TLS_ECDHE_RSA_WITH_AES_128_CBC_SHA
 TLS_ECDHE_ECDSA_WITH_AES_128_CBC_SHA
 TLS_ECDHE_RSA_WITH_AES_256_CBC_SHA
 TLS_ECDHE_ECDSA_WITH_AES_256_CBC_SHA
 TLS_RSA_WITH_RC4_128_SHA
 TLS_RSA_WITH_AES_128_CBC_SHA
 TLS_RSA_WITH_AES_256_CBC_SHA
 TLS_ECDHE_RSA_WITH_3DES_EDE_CBC_SHA
 TLS_RSA_WITH_3DES_EDE_CBC_SHA
Extensions (6):
 server_name
 status_request
 elliptic_curves
 ec_point_formats
 signature_algorithms
 renegotiation_info

```

**(b) Firefox 31**
```
Ciphersuites (23):
 TLS_ECDHE_ECDSA_WITH_AES_128_GCM_SHA256
 TLS_ECDHE_RSA_WITH_AES_128_GCM_SHA256
 TLS_ECDHE_ECDSA_WITH_AES_256_CBC_SHA
 TLS_ECDHE_ECDSA_WITH_AES_128_CBC_SHA
 TLS_ECDHE_RSA_WITH_AES_128_CBC_SHA
 TLS_ECDHE_RSA_WITH_AES_256_CBC_SHA
 TLS_ECDHE_RSA_WITH_3DES_EDE_CBC_SHA
 TLS_ECDHE_ECDSA_WITH_RC4_128_SHA
 TLS_ECDHE_RSA_WITH_RC4_128_SHA
 TLS_DHE_RSA_WITH_AES_128_CBC_SHA
 TLS_DHE_DSS_WITH_AES_128_CBC_SHA
 TLS_DHE_RSA_WITH_CAMELLIA_128_CBC_SHA
 TLS_DHE_RSA_WITH_AES_256_CBC_SHA
 TLS_DHE_DSS_WITH_AES_256_CBC_SHA
 TLS_DHE_RSA_WITH_CAMELLIA_256_CBC_SHA
 TLS_DHE_RSA_WITH_3DES_EDE_CBC_SHA
 TLS_RSA_WITH_AES_128_CBC_SHA
 TLS_RSA_WITH_CAMELLIA_128_CBC_SHA
 TLS_RSA_WITH_AES_256_CBC_SHA
 TLS_RSA_WITH_CAMELLIA_256_CBC_SHA
 TLS_RSA_WITH_3DES_EDE_CBC_SHA
 TLS_RSA_WITH_RC4_128_SHA
 TLS_RSA_WITH_RC4_128_MD5
Extensions (8):
 server_name
 renegotiation_info
 elliptic_curves
 ec_point_formats
 SessionTicket TLS
 next_protocol_negotiation
 status_request
 signature_algorithms

```

**(c) Chrome 40**
```
Ciphersuites (18):
 TLS_ECDHE_ECDSA_WITH_AES_128_GCM_SHA256
 TLS_ECDHE_RSA_WITH_AES_128_GCM_SHA256
 TLS_DHE_RSA_WITH_AES_128_GCM_SHA256
 TLS_ECDHE_ECDSA_WITH_AES_256_CBC_SHA
 TLS_ECDHE_ECDSA_WITH_AES_128_CBC_SHA
 TLS_ECDHE_RSA_WITH_AES_128_CBC_SHA
 TLS_ECDHE_RSA_WITH_AES_256_CBC_SHA
 TLS_ECDHE_ECDSA_WITH_RC4_128_SHA
 TLS_ECDHE_RSA_WITH_RC4_128_SHA
 TLS_DHE_RSA_WITH_AES_128_CBC_SHA
 TLS_DHE_DSS_WITH_AES_128_CBC_SHA
 TLS_DHE_RSA_WITH_AES_256_CBC_SHA
 TLS_RSA_WITH_AES_128_GCM_SHA256
 TLS_RSA_WITH_AES_128_CBC_SHA
 TLS_RSA_WITH_AES_256_CBC_SHA
 TLS_RSA_WITH_3DES_EDE_CBC_SHA
 TLS_RSA_WITH_RC4_128_SHA
 TLS_RSA_WITH_RC4_128_MD5
Extensions (10):
 server_name
 renegotiation_info
 elliptic_curves
 ec_point_formats
 SessionTicket TLS
 next_protocol_negotiation
 Application Layer Protocol Negotiation
 Channel ID
 status_request
 signature_algorithms

```

**Fig. 12. Selected differences in ClientHello messages in three different TLS implementations. Even though the contents of application**
data records are hidden by encryption, the plaintext headers of TLS reveal information about the implementation. This figure illustrates the need to disguise the TLS fingerprint so that it is not easily identified as pertaining to a circumvention tool.


-----