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Vulnerabilities in the Kerberos version 4 protocol
Date: 2003-03-16

Author : hack4life@hushmail.com


Several cryptographic vulnerabilities exist in the basic Kerberos
Version 4 protocol that could allow an attacker to impersonate any
user in a Kerberos realm and gain any privilege authorized through
that Kerberos realm. Knowledge of the key shared between two realms
for Kerberos 4 cross-realm authentication or the ability to create
arbitrary principals in a realm is sufficient to print any ticket in
the realm. As an example, knowing krbtgt.ZONE.MIT.EDU@ATHENA.MIT.EDU
is sufficient to print an Athena TGT for any Athena realm client.
Additional vulnerabilities in a MIT extension to use tripple DES
keys for Kerberos 4 tickets may allow attackers who can passively
observer the network to construct tickets for some users if certain
alignment constraints are met.

The Kerberos 5 protocol is not vulnerable to this issue. However,
implementations that implement both Kerberos 4 and Kerberos 5 tend to
use the same keys for both protocols. As a result, the Kerberos 4
vulnerabilities can be used to compromise Kerberos 5 services with
these implementations.

Brief Problem Description

Kerberos version 4 tickets include neither a cryptographic hash of the
encrypted data, random padding, nor a random initial vector. As such,
if an attacker can cause the right text to be encrypted in a Kerberos
service key, then the attacker can fabricate a ticket. Normally an
attacker does not control much of the text in the ticket so this
cryptographic weakness is hard to exploit.

The initial portion of a Kerberos 4 ticket is a one-byte flags field (either 0
or 1) followed by the client name. Since all of this initial text is
constant, the beginning of a ticket for a given client/service will be
the same. An attacker thus knows the encryption of the initial
plaintext in the service key. If an attacker can control client
principals whose names he chooses, then he can get the encryption of
these plaintext values in the service key. An attached paper
details how to choose principal names in order to encrypt arbitrary
plaintext and how to use this ability to construct tickets for both
Kerberos 4 and Kerberos 5.

As a result of concerns about single DES weaknesses, MIT implemented
support for Kerberos 4 tickets encrypted in triple DES service keys.
This support shares all the cryptographic weaknesses of single DES
Kerberos 4. In addition, since it uses CBC mode rather than PCBC
mode, it introduces new weaknesses not found in other Kerberos 4
implementations. When certain alignment constraints are met, it is
possible to splice two tickets together, allowing an attacker to get a
ticket with a known session key for a client without knowing that
client's long term key. This attack does require sniffing a ticket
for that client.

We do not believe the password changing service is vulnerable to the
single DES attacks as the KDC will never issue password changing
tickets in an appl request. It is probably vulnerable to the triple
DES splicing attacks.

Specific Vulnerabilities

1) ECB Oracle for Single DES

By controlling principals of an attackers choice, an attacker can
encrypt arbitrary plaintext in a single DES service key.

2) ECB Oracle for Triple DES

By controlling principals of an an attacker's choice, an attacker
can encrypt arbitrary plaintext in a triple DES service key.

3) PCBC First Block

It turns out that being able to encrypt arbitrary plaintext is not
quite enough to construct a ticket for a single DES service key.
You also need to be able to construct the first block of the
ticket; you don't know what plaintext to use because the IV for
the first block is the long-term service key. However since the
only thing in the first block of the ticket is the first seven
bytes of the client, controlling a principal with the same first
seven bytes as the principal being attacked is sufficient to get
the first block. As a practical matter, principals whose
principal and instance components fit within six bytes (including
trailing nulls) may be harder to attack. The attached paper
discusses mechanisms for mounting such an attack.

4) Cross Realm

If realms A and B share a cross-realm key and the attacker knows
that key or can get arbitrary plaintext encrypted in that key,
then the attacker may get A to issue tickets for any principal
claiming to be in realm B and vice versa. This is sufficient to
meet conditions of vulnerabilities 1 and 2 above and to encrypt
arbitrary plaintext in the service keys of realm A and B.

5) Kerberos 4 Ticket Printing

The conditions of 2 above are sufficient to print arbitrary
tickets in a triple DES service key. The conditions of 1 and 3
are sufficient to print any ticket in a single DES service key.

6) Kerberos 5 Ticket Printing

The conditions of 1 above are sufficient to construct a
des-cbc-md4 or des-cbc-md5 Kerberos 5 ticket if the KDC uses the
same DES key for v4 and v5. While the Kerberos 5 ticket does have
a confounder and checksum, the checksum is not keyed and thus the
confounder and checksum can be fabricated by an attacker. We
believe that des-cbc-crc is safe unless you can contain a
ciphertext block and a corresponding plaintext block; the paper
discusses situations where this is possible. However most
Kerberos implementations will allow des-cbc-md5 to be used even if
des-cbc-crc is normally used. We are not aware of any
vulnerabilities in des3-hmac-sha1-kd or rc4-hmac-md5.

7) Ticket Splicing Attack

A Kerberos 4 ticket contains an eight-byte session key. If client
principal names are chosen carefully then this session key will
line up with a DES block boundary. For triple DES service keys
this creates an opportunity for an attack. Consider the case
where an attacker has obtained a ticket t1 with a known session
key K and has sniffed a ticket t2 with unknown session key for the
same service. The attacker can create a new valid ticket t2' by
replacing the part of t2 starting with the session key block with
the session key from t1. This new ticket will have a session key
K XOR-ed with the ciphertext blocks proceeding the session key in
t1 and t2. In other words, if triple DES service keys are used,
client principals with the wrong name lengths are inherently
vulnerable to sniffing.

8) Realm Hopping

Kerberos 4 does not normally support multi-hop cross-realm
authentication. However cross-realm tickets are just normal
service keys; points 1,2 and 3 are sufficient to satisfy the
conditions of point 4 for a service key. That is, an attacker
can hop through realms, exploiting these vulnerabilities against
any realm that is in the transitive closure of the initial realm.
Anyone who shares keys with ATHENA.MIT.EDU now trusts

9) Krb 524 Does Not Help

Traditionally realms desiring higher security but still wishing to
have some Kerberos 4 services have disabled KDC support for V4 and
used krb524d to issue only the services that are needed. These
vulnerabilities work as well against any service key that the krb524d
knows as they do against service keys in a v4 KDC. Of course a
fabricated krb5 ticket can be converted to Kerberos 4 using krb524d.

Potential Solutions

1) V4 Cross Realm Considered Harmful

Kerberos implementations should gain an option to
disable Kerberos 4 cross-realm authentication both in the KDC and
in any implementations of the krb524 protocol. This configuration
should be the default.

2) Application Migration

Application vendors and sites should migrate from Kerberos version 4
to Kerberos version 5. The OpenAFS community has introduced features
that allow Kerberos 5 to be used for AFS in OpenAFS 1.2.8. Patches
are available to add Kerberos 5 support to OpenSSH. Several other
implementations of the SSH protocol also support Kerberos 5.
Applications such as IMAP, POP and LDAP already support Kerberos 5.

3) TGT Key Separation

One motivation for the V4 triple DES support is that if a single
DES key exists for the TGT principal then an attacker can attack
that key both for v4 and v5 tickets. Kerberos
implementations should gain support for a DES TGT key that is used
for v4 requests but not v5 requests.

4) Remove Triple DES Kerberos 4 Support

The cut and paste attack is a critical failure in MIT's attempt at
Kerberos 4 Triple DES. Even without cross-realm authentication,
this can be exploited in real-world situations. As such the
support for 3DES service keys should be disabled.

Obtaining Patches

The following Kerberos implementations have provided statements on how
vendors should obtain patches:

MIT: Patches are available for the 1.2.x and 1.3 versions of MIT
Kerberos 5. Contact Sam Hartman <hartmans@mit.edu> or Tom Yu
<tlyu@mit.edu> for patches.


This description was written by Sam hartman. The attached paper
detailing the cryptographic details of the attack was written by Tom
Yu. The exploit was written by Ken Raeburn and Sam Hartman. All
parties were involved in discovering the vulnerabilities.

*) Paper by Tom on weaknesses

- ----------------------------------------


The Kerberos 4 protocol is vulnerable to a number of adaptive
chosen-plaintext attacks. Some of these attacks may be used,
indirectly, against Kerberos 5 as well, and an attack exists that can
generate arbitrary tickets in any realm that either directly or
transitively shares a key with a realm controlled by an adversary.

These attacks have their basis in the ability of an adversary to
create an Electronic Code Book (ECB) oracle for a block cipher with a
fixed unknown key, given:

* the ability to choose a single block of a larger plaintext for a
victim to encrypt with that key using the block cipher in a chaining

* the output ciphertext block of the chaining-mode encryption
corresponding to the chosen plaintext

* prior knowledge of the feedback block that will be XORed with the
chosen plaintext block prior to ECB-encryption

The existence of this oracle permits the attacker, without knowledge
of the key, to construct a ciphertext that will decrypt with a chained
block cipher to any desired plaintext (with the possible exception of
the first block). This constructed ciphertext can be the entirety of
an arbitrary Kerberos 4 ticket. Kerberos 4 is vulnerable to this
attack due to the lack of random confounders, the lack of random
initialization vectors, and the lack of cryptographically strong
integrity checking in its use of block ciphers. Kerberos 5 is
somewhat less vulnerable, but a related weakness permits a
cross-protocol attack from Kerberos 4 to Kerberos 5 when single-DES is
used to encrypt tickets. In addition, these vulnerabilities render
any realm that either directly or transitively shares a key with an
adversary-controlled realm vulnerable to compromise.

Definitions and notation

[ heavily adapted from B. Schneier, _Applied Cryptography_, 2nd ed.,
John Wiley & Sons, Inc., 1996, chapter 9 ]

A ^ B = A XOR B

C[n] = nth ciphertext block
P[n] = nth plaintext block
E(k, x) = x ECB-encrypted with key k
D(k, x) = x ECB-decrypted with key k

Cipher Block Chaining (CBC) mode is used in the implementation of
triple-DES in krb4 that is used in the MIT krb5 release:

C[n] = E(k, P[n] ^ C[n-1])
P[n] = D(k, C[n]) ^ C[n-1]

Propagating Cipher Block Chaining (PCBC) mode is used for single-DES
encryption in krb4:

C[n] = E(k, P[n] ^ C[n-1] ^ P[n-1])
P[n] = D(k, C[n]) ^ C[n-1] ^ P[n-1]

It is useful to generalize these block cipher chaining modes by
considering the block that is to be ECB-encrypted as being a plaintext
block XORed with a feedback block:

F[n] = nth feedback block (not transmitted)
C[n] = E(k, P[n] ^ F[n])
P[n] = D(k, C[n]) ^ F[n]

F[0] = initialization vector (IV)

for CBC:

F[n] = C[n-1]

for PCBC:

F[n] = C[n-1] ^ P[n-1]

There are other, more obscure modes, e.g. Block Chaining (BC) mode:

F[n] = F[n-1] ^ C[n-1]

but this document will not discuss them further, since the krb4
protocol does not use them.

General ECB oracle attack on chained block ciphers

An ECB oracle for a block cipher can be constructed from a
block-chaining mode of that cipher if an attacker has access to the
ciphertext block corresponding to a chosen plaintext block of a larger
plaintext, under certain conditions.

To learn E(k, x) for arbitrary x, note that

C[n] = E(k, P[n] ^ F[n])

in the generalized form for chained block ciphers. Let

x = P[n] ^ F[n] ,

and assume that the attacker has prior knowledge of the block F[n]
that will be used when the victim encrypts the complete plaintext
containing the chosen plaintext block P[n]. The attacker may then
manipulate the chosen plaintext block P[n] to produce

C[n] = E(k, x)

by choosing

P[n] = x ^ F[n] ,

since the XOR operation is its own inverse. The attacker now has an
oracle that can ECB-encrypt arbitrary plaintext using the key k,
without knowing the key. The complete set of requirements for the
attacker to construct this ECB oracle is:

* the ability to choose a plaintext block P[n] as part of a (possibly)
larger plaintext to be chaining-mode encrypted with the fixed key k

* prior knowledge of the feedback block F[n] that will be XORed with
P[n] when this chaining-mode encryption occurs

* knowledge of the ciphertext block C[n] that results from the
chaining-mode encryption

In a well-designed cryptosystem, an attacker should have significant
difficulty assembling such an oracle, since it should be difficult to
determine F[n] ahead of time. If F[n] is known to be constant between
two instances of the chaining-mode encryption, an attacker may rather
easily construct the ECB oracle.

As an example of how to mount this attack, assume that the attacker
has access to the output of a black box that produces ciphertext that
is encrypted in a chaining mode with a fixed key k. Additionally,
assume that the attacker can vary a constrained subset of inputs to
the black box such that the only resulting changes in the ciphertext,
up to and including C[n], are in C[n] itself, and that the feedback
block F[n] associated with C[n] does not change. The ciphertext
following C[n] is not important. If the attacker can manipulate the
constrained subset of the black box's inputs in a way that completely
controls the contents of the plaintext block P[n] that will encrypt to
C[n], then the attacker has an ECB oracle for the fixed key k.

Note that the attacker does _not_ need to control the entire plaintext
that the black box encrypts; it is sufficient to manipulate one
aligned block of plaintext. Also, note that the condition of fixed
output ciphertext up through C[n] is slightly less restrictive in the
general case; the attacker only needs to ensure that F[n] does not
vary with manipulation of the constrained subset of inputs to the
black box, though this usually implies that the output ciphertext up
to C[n] is also fixed.

Using an ECB oracle to construct arbitrary ciphertext

To construct a complete ciphertext that a chained block cipher will
decrypt to a desired plaintext, it is necessary to proceed one block
at a time, taking into account the feedback block each time. Since

C[n] = E(k, P[n] ^ F[n]) ,

it is only necessary to take the desired plaintext block P[n], XOR it
with the feedback block F[n], and encrypt it using the ECB oracle.
The construction of C[0] for a desired P[0] using the ECB oracle is
only easily possible if the initialization vector F[0] that will be
used for decryption is known. If F[0] is not known, it will be
necessary to manipulate P[0] to be encrypted in the chaining mode by k
using the unknown value of F[0], which is not in general possible. In
general, the attacker will know F[n] for non-zero values of n because
F[n] will depend solely on the previously constructed blocks of
ciphertext (or, additionally, the previous plaintext block in the case
of PCBC).

In the specific case of single-DES used in PCBC mode in krb4, F[0] is
the key k, and is not known to an attacker, even if the ECB oracle can
be obtained. To produce a C[0] that decrypts to a desired value, the
attacker must be able to manipulate the P[0] that will be encrypted by

In the case of triple-DES used in CBC mode used in recent krb4
implementations in the krb5 sources, F[0] is an all-zeros IV, so the
ECB oracle can be used to produce a C[0] that decrypts to an arbitrary
P[0] even if P[0] cannot be directly manipulated to be encrypted with
k in CBC mode by the attacker.

Building an ECB oracle using krb4 tickets

A krb4 ticket is encrypted in the long-term key of the service. The
plaintext contents of the ticket are:

unsigned char flags namely, HOST_BYTE_ORDER
string pname client's name
string pinstance client's instance
string prealm client's realm
4 bytes paddress client's address
8 bytes session session key
1 byte life ticket lifetime
4 bytes time_sec KDC timestamp
string sname service's name
string sinstance service's instance
<=7 bytes null null pad to 8 byte multiple

The fields labeled "string" are NUL-terminated ASCII strings, each
limited to 40 characters (including the terminal NUL).

The following example of ECB oracle construction assumes that the
attacker varies P[1] to obtain the ECB-encryption of a desired
plaintext in C[1]. It may be generalized to use manipulation of an
arbitrary P[n], though.

Assume that the attacker controls a realm "HAX0R.EXAMPLE", and the
attacked realm is "VICTIM.EXAMPLE", and that they share a key. Since
the realms "HAX0R.EXAMPLE" and "VICTIM.EXAMPLE" share a key, the
attacker knows the key for the principal


(which has the same key as


in krb4). The attacker can fabricate a cross-realm ticket for the
client principal


where "a234567" is a constant string, and "XXXXXXXX" is chosen to
produce the ECB encryption of an arbitrary plaintext by the key to be
attacked. The service principal of this magic ticket will of course
be for the service principal


since it is a cross-realm TGT.

The attacker then uses the cross-realm ticket to obtain a service
ticket for the key that will be attacked, for example,


(for maximum damage). The attacked KDC (which acts as the black box
assumed in the section "General ECB oracle attack") encrypts the
client principal's name, instance, and realm from the cross-realm TGT
with the attacked service's key, thus providing ciphertext
corresponding to a partially known plaintext. Note that the session
key, timestamp, etc. may not be known to the attacker before
encryption, but they do not affect the ciphertext blocks of interest.

The fixed string "a234567", appended to the one-byte flags field,
becomes the first plaintext block P[0]. This means that chosen
plaintext string "XXXXXXXX" is the entirety of the second plaintext
block P[1]. C[0] will be a fixed value, given the fixed string, since
the flags don't change, the first 7 characters of the principal name
don't change, the IV and key don't change, and there is no prepended
confounder (unlike in krb5 uses of block-chaining modes). Likewise,
F[1] won't change, since it is based on fixed values (C[0], which is
additionally XORed with P[0] in PCBC mode). The attacker merely has
to change the variable string "XXXXXXXX", which is P[1], to be the
result of F[1] XORed with the desired plaintext to be ECB-encrypted.
The C[1] in the resulting ticket will then be the ECB encryption of
the desired plaintext. The attacker can obtain the initial fixed
value of F[1] by performing one pass with the chosen P[0] and
discarding C[1].

If any P[1] obtained by XORing F[1] with the desired plaintext
contains one NUL character, the value of P[1] will be split between
the "pname" and "pinstance" fields, rather than being completely
contained within the "pname" field. This is still acceptable. If
there are two or more NUL characters in the resulting P[1], the KDC
will probably not create a useful ticket, so the previously fixed
string P[0] must be permuted to obtain a F[1] value that does not
result in an unacceptable number of NUL characters in P[1]. It may
also be possible to permute P[0] such that P[1] will not need to
contain any "suspicious" characters, e.g. non-alphanumerics, in order
to obtain a desired C[1].

Note that almost none of this attack is specific to the particular
block-chaining mode (CBC or PCBC), provided that the mode is of the
general form described under "Definitions".

Note also that it is not necessary to control a cross-realm key with
the realm "VICTIM.EXAMPLE"; it is merely sufficient to be able to
create or control a large number of principals in "VICTIM.EXAMPLE" and
to know their keys. Having control of a cross-realm key merely makes
the attack much easier.

Even that level of control may not be required, if the attacker has
control of a key for a principal whose name is of the correct length
to cause the "paddress" field to begin on a block boundary. Then,
assuming that the attacker can receive arbitrary UDP packets
originating from the attacked KDC and that are addressed to arbitrary
IP addresses, the attacker can effectively vary the "paddress" field,
by sending from arbitrary source IP addresses, in order to obtain
control over 4 bytes of plaintext. Note that this increases the work
factor to O(2**32) per ciphertext block, since the immediately
following 4 bytes are the first half of the (presumed) random session
key, which the attacker can obtain only after the KDC issues the
ticket. Obviously, controlling smaller amounts of address space will
result in a proportionally increased work factor, but it can still be
significantly below O(2**56). The downside to this attack is that the
large volume of ticket requests required may be quite noticeable to a
KDC administrator.

Most of the the "life" field and some of the "time_sec" field are also
under the control of an attacker, and can be used as chosen plaintext
with which to mount the ECB oracle attack, though this method may be
somewhat more difficult, and suffers from the same drawbacks as the
address space attack.

Constructing an arbitrary krb4 ticket using the ECB oracle

Given the ECB oracle, it is possible to construct an arbitrary krb4
ticket, subject to certain conditions. The attacker must usually
build up the ciphertext of the fabricated ticket one block at a time,
since the feedback blocks are required; there may be optimizations
that permit multiple ciphertext blocks to be produced and spliced in
at once, but the length constraints inherent in krb4 principal names,
in addition to the NUL-terminated string condition, limit these
optimizations somewhat.

An arbitrary block of ciphertext C[n] corresponding to a chosen
plaintext P[n] may be produced as described above in "General ECB
oracle attack". Production of the first ciphertext block, C[0], may
be difficult if F[0] is not known. This is the case in krb4
single-DES in PCBC mode, since the key is used as the initialization
vector F[0]. In the krb4 usage of triple-DES CBC mode, the IV is all
zeros, so it is not necessary to force encryption of a chosen P[0] in
order to obtain a desired C[0].

Since, in the krb4 usage of single-DES PCBC, C[0] will be constant for
a constant P[0], it is sufficient to somehow obtain a C[0]
corresponding to the desired P[0]. This may be trivially done by
controling a realm with a shared key, since the C[0] may be obtained
directly in that case. Consider an attacker who wants to fabricate a
ticket for

"someluser@VICTIM.EXAMPLE" .

The attacker may simply synthesize a cross-realm ticket for the client

"somelu@HAX0R.EXAMPLE" ,

since the resulting P[0],


ends immediately prior to the terminating NUL of the "somelu"
principal name string. Upon using the cross-realm ticket to obtain a
service ticket in the "VICTIM.EXAMPLE" realm, the first block of the
ticket will be exactly the C[0] required to synthesize the remainder
of the ticket, even though F[0] is unknown. Note that fabricating a
ticket in this manner for a client whose total length of principal
name and instance (including terminating NUL characters) is less than
7 bytes is probably impossible unless the attacker's realm name and
the victim's realm name share an initial substring. This constraint
is a consequence of the requirement that the client realm name in the
ticket must match the issuing realm of the ticket.

Without access to a realm with a shared key, it suffices to sniff a
C[0] corresponding to a desired P[0] off the network. This may be
difficult if the client principal(s) being attacked are careful not to
obtain or use krb4 tickets for authentication. There are other ways
to obtain the desired initial ciphertext block C[0], as described in
later sections.

Cross-protocol attack from krb4 to krb5

If a service is authenticated with krb5, accepts tickets encrypted
with single-DES, and does not necessarily accept krb4 authentication,
it is still possible to mount an attack against it using the krb4 ECB
oracle if the KDC for the service's realm can issue krb4 tickets for
the service. The format of a krb5 single-DES plaintext immediately
prior to encryption is

concat(confounder, checksum, data, pad) .

The confounder is one block of random bytes. The checksum may be
CRC-32, MD4, or MD5. None of these checksums is keyed, which is the
crucial feature that enables this attack. For des-cbc-md4 and
des-cbc-md5, the IV is all zeros. For des-cbc-crc, the IV is the key.
The checksum is computed over the entire concatenated plaintext, with
the checksum field zeroed out. This frustrates ciphertext generation
via chosen-plaintext manipulation if the attacker has no knowledge of
the confounder; an adversary will find it very difficult, if not
impossible, to create an ECB oracle based on manipulating krb5

For des-cbc-md4 and des-cbc-md5, an attacker can simply fabricate any
desired confounder using the ECB oracle, since the IV is known and
constant. For des-cbc-crc, the attack is somewhat more difficult,
since the IV is the key and not known to the attacker. For this
particular attack, it is necessary to obtain the ciphertext block C[0]
corresponding to the desired confounder, possibly by the same method
that can be used for the C[0] attack on PCBC krb4, i.e. by forcing the
KDC to encrypt an intial P[0] with the attacked key (which also
requires control of a realm with a shared key).

An easier method for obtaining C[0] for attacking krb5 des-cbc-crc is
to sniff, or otherwise acquire, any krb4 ticket encrypted in the
attacked key, provided that the first block of plaintext is known.
The first ciphertext block of both the krb5 des-cbc-crc encryption and
the krb4 PCBC encryption of the same plaintext will be identical, as
they both use the same key and an IV equal to the key.

Incidentally, if krb524d is running on the attacked KDC, the
cross-protocol attack from krb4 to krb5 may be extended back into an
attack on single-DES krb4, circumventing the problems of constructing
C[0] for the krb4 ticket, especially the difficulties inherent in
fabricating a ticket for a client with a short principal name. The
attacker acquires C[0] of a krb4 ticket encrypted for the target
service, provided that the corresponding P[0] is known. Note that
this is trivial if the attacker knows _any_ client principal key in
the attacked realm. The attacker can then use this C[0] as the
encrypted confounder for a krb5 des-cbc-crc ticket, and proceed to
construct a krb5 ticket for that service, complete with correct
checksum. Then, the attacker merely submits the fabricated krb5
ticket to krb524d, which converts it into a valid krb4 ticket.

This weakness of the single-DES krb5 cryptosystems results largely
from the practice of encrypting plaintext checksums. Without knowing
the key, an attacker can still fabricate ciphertext that decrypts and
appears to have valid integrity. [ S. M. Bellovin, personal
communication, 1999 ] This would not be possible if a keyed checksum
or a HMAC were used. Note, however, that checksumming of the random
confounder along with the data prevents several types of
chosen-plaintext cut-and-paste attacks.

Attack on transitive closure of trust

Normally, cross-realm trust in krb4 is not transitive, because the
code in the KDC implementation forbids realm-hopping by use of
cross-realm tickets. This can be circumvented by an adversary,
though. The nature of cross-realm trust in krb4 means that any remote
realm, if it can synthesize tickets for services in the local realm,
can synthesize tickets for services in any other realm sharing keys
with the local realm. Consider a local target realm "VICTIM.EXAMPLE",
which shares keys with an attacker's realm "HAX0R.EXAMPLE" and with
another target realm "OWNZ0RED.EXAMPLE". The attacker can not only
synthesize tickets for clients in "VICTIM.EXAMPLE" for services in
"VICTIM.EXAMPLE", but can also synthesize a ticket for an arbitrary
client in the "VICTIM.EXAMPLE" realm for the service


This ability to synthesize arbitrary cross-realm tickets between
"VICTIM.EXAMPLE" and "OWNZ0RED.EXAMPLE" means that the attacker can
synthesize a ticket for a client in "OWNZ0RED.EXAMPLE" for any service
in "OWNZ0RED.EXAMPLE" by using the ECB oracle attack against the
service key. The attacker may recurse this attack indefinitely,
compromising the entire web of trust. Combined with the krb4-to-krb5
cross-protocol attack, the results may be devastating.


The Kerberos 4 protocol has weaknesses in its uses of cryptography
that permit an attacker to mount an adaptive chosen-plaintext attack
to obtain an ECB oracle. This oracle can then be used to fabricate
arbitrary tickets in any realm that shares a key with one under the
attacker's control. This attack has some limitations, but the
combination of cross-protocol attack to the Kerberos 5 protocol with
the related cross-protocol attack back to Kerberos 4 permits the
attacker to circumvent most of these limitations. Additionally, these
fabricated tickets permit the attacker to extend the attack to all
transitively trusted realms, in spite of the normal lack of
cross-realm trust transitivity in the Kerberos 4 protocol. Combined,
these attacks can have a profound detrimental effect on a set of
realms that trust each other.


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