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CHAPTER 19 ................................................ 460
Public-Key Algorithms .............................. 460
BACKGROUND ....................................... 460
Security Public-Key Algorithms ..................... 460
KNAPSACK ALGORITHMS ..................... 461
Superincreasing Knapsacks .......................... 462
Figure 19.1 Encryption with knapsacks. ......... 462
Creating the Public Key from the Prioate ....... 463
Encryption ..................................................... 463
Decryption ..................................................... 464
Practical Implementations ............................. 464
Security of Knapsacks ................................... 464
Knapsack Variants ........................................ 464
Patents .......................................................... 465
RSA ........................................................... 465
Table 19.1 Foreign Merkle-Hellman .............. 465
Table 19.2 RSA Encryption .......................... 467
Public Key: ................................................ 467
Private Key: .............................................. 467
Encrypting: ................................................ 467
Decrypting: ............................................... 467
RSA in Hardware .......................................... 468
Speed RSA ................................................... 468
Software Speedups ....................................... 468
Table 19.3 Existing RSA Chips .................... 468
Table 19.4 RSA Speeds for Different ............ 469
Security RSA ................................................. 469
Chosen Ciphertext Attack against RSA ........ 470
Common Modulus Attack on RSA ................. 471
Low Encryption Exponent Attack against ....... 471
Low Decryption Exponent Attack against ....... 472
Lessons Learned ........................................... 472
Attack on Encrypting and Signing with ........... 472
Standards ...................................................... 473
Patents .......................................................... 473
POHLIG-HELLMAN ................................. 473
Patents .......................................................... 473
RABIN ...................................................... 474
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Williams ......................................................... 474
ELGAMAL ................................................. 475
ElGamal Signatures ...................................... 475
Table 19.5 ElGamal Signatures ................... 476
ElGamal Encryption ...................................... 477
Speed ............................................................ 477
Table 19.6 ElGamal Encryption ................... 477
Patents .......................................................... 478
MCELIECE ............................................... 478
Table 19.7 ElGamal Speeds for Different ..... 478
Other Algorithms Based on Linear ................. 479
ELLIPTIC CRYPTOSYSTEMS ................ 479
LUC ........................................................... 480
FINITE AUTOMATON PUBLIC-KEY......... 481
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19
CHAPTER


Public-Key Algorithms


19.1 BACKGROUND
The concept of public-key cryptography was invented by Whitfield Diffie and Mar-
tin Hellman, and independently by Ralph Merkle. Their contribution to cryptogra-
phy was the notion that keys could come in pairs-an encryption key and a
decryption key-and that it could be infeasible to generate one key from the other
(see Section 2.5). Diffie and Hellman first presented this concept at the 1976
National Computer Conference [495]; a few months later, their seminal paper “New
Directions in Cryptography” was published (4961. (Due to a glacial publishing pro-
cess, Merkle™s first contribution to the field didn™t appear until 1978 [1064].)
Since 1976, numerous public-key cryptography algorithms have been proposed.
Many of these are insecure. Of those still considered secure, many are impractical.
Either they have too large a key or the ciphertext is much larger than the plaintext.
Only a few algorithms are both secure and practical. These algorithms are gener-
ally based on one of the hard problems discussed in Section 11.2. Of these secure and
practical public-key algorithms, some are only suitable for key distribution, Others
are suitable for encryption (and by extension for key distribution). Still others are
only useful for digital signatures. Only three algorithms work well for both encryp-
tion and digital signatures: RSA, ElGamal, and Rabin. All of these algorithms are
slow. They encrypt and decrypt data much more slowly than symmetric algorithms;
usually that™s too slow to support bulk data encryption.
Hybrid cryptosystems (see Section 2.5) speed things up: A symmetric algorithm
with a random session key is used to encrypt the message, and a public-key algo-
rithm is used to encrypt the random session key.
of Public-Key
Security Algorithms
Since a cryptanalyst has access to the public key, he can always choose any mes-
sage to encrypt. This means that a cryptanalyst, given C = E,(P), can guess the value
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19 Public-Key Algorithms
CHAPTER



of P and easily check his guess. This is a serious problem if the number of possible
plaintext messages is small enough to allow exhaustive search, but can be solved by
padding messages with a string of random bits. This makes identical plaintext mes-
sages encrypt to different ciphertext messages. (For more about this concept, see
Section 23.15.)
This is especially important if a public-key algorithm is used to encrypt a session
key. Eve can generate a database of all possible session keys encrypted with Bob™s
public key. Sure, this requires a large amount of time and memory, but for a 40-bit
exportable key or a 56-bit DES key, it™s a whole lot less time and memory than
breaking Bob™s public key. Once Eve has generated the database, she will have his
key and can read his mail at will.
Public-key algorithms are designed to resist chosen-plaintext attacks; their secu-
rity is based both on the difficulty of deducing the secret key from the public key
and the difficulty of deducing the plaintext from the ciphertext. However, most
public-key algorithms are particularly susceptible to a chosen-ciphertext attack (see
Section 1.1).
In systems where the digital signature operation is the inverse of the encryption
operation, this attack is impossible to prevent unless different keys are used for
encryption and signatures.
Consequently, it is important to look at the whole system and not just at the indi-
vidual parts. Good public-key protocols are designed so that the various parties can™t
decrypt arbitrary messages generated by other parties-the proof-of-identity proto-
cols are a good example (see Section 5.2).


19.2 KNAPSACK ALGORITHMS
The first algorithm for generalized public-key encryption was the knapsack algo-
rithm developed by Ralph Merkle and Martin Hellman [713,1074]. It could only be
used for encryption, although Adi Shamir later adapted the system for digital signa-
tures [1413]. Knapsack algorithms get their security from the knapsack problem, an
NP-complete problem. Although this algorithm was later found to be insecure, it is
worth examining because it demonstrates how an NP-complete problem can be
used for public-key cryptography.
The knapsack problem is a simple one. Given a pile of items, each with different
weights, is it possible to put some of those items into a knapsack so that the knap-
sack weighs a given amount? More formally: Given a set of values Ml, Ml, . . . , M,,
and a sum S, compute the values of bi such that
S=b,M,+b2M,+...+b,M,
The values of bi can be either zero or one. A one indicates that the item is in the
knapsack; a zero indicates that it isn™t.
For example, the items might have weights of 1, 5, 6, 11, 14, and 20. You could
pack a knapsack that weighs 22; use weights 5,6, and 11. You could not pack a knap-
sack that weighs 24. In general, the time required to solve this problem seems to
grow exponentially with the number of items in the pile.
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19.2 Knapsack Algorithms


The idea behind the Merkle-Hellman knapsack algorithm is to encode a message
as a solution to a series of knapsack problems. A block of plaintext equal in length
to the number of items in the pile would select the items in the knapsack (plaintext
bits corresponding to the b values), and the ciphertext would be the resulting sum.
Figure 19.1 shows a plaintext encrypted with a sample knapsack problem.
The trick is that there are actually two different knapsack problems, one solvable
in linear time and the other believed not to be. The easy knapsack can be modified
to create the hard knapsack. The public key is the hard knapsack, which can easily
be used to encrypt but cannot be used to decrypt messages. The private key is the
easy knapsack, which gives an easy way to decrypt messages. People who don™t
know the private key are forced to try to solve the hard knapsack problem.
Superincreasing Knapsacks
What is the easy knapsack problem? If the list of weights is a superincreasing
sequence, then the resulting knapsack problem is easy to solve. A superincreasing
sequence is a sequence in which every term is greater than the sum of all the previ-
ous terms. For example, (1,3,6,13,27,52} is a superincreasing sequence, but (1,3,4,9,
l&2.5] is not.
The solution to a superincreasing knapsack is easy to find. Take the total weight
and compare it with the largest number in the sequence. If the total weight is less
than the number, then it is not in the knapsack. If the total weight is greater than or
equal to the number, then it is in the knapsack. Reduce the weight of the knapsack
by the value and move to the next largest number in the sequence. Repeat until fin-
ished. If the total weight has been brought to zero, then there is a solution. If the
total weight has not, there isn™t.
For example, consider a total knapsack weight of 70 and a sequence of weights of
(2,3,6,13,27,52]. The largest weight, 52, is less than 70, so 52 is in the knapsack. Sub-
tracting 52 from 70 leaves 18. The next weight, 27, is greater than 18, so 27 is not in
the knapsack. The next weight, 13, is less than 18, so 13 is in the knapsack. Sub-
tracting 13 from 18 leaves 5. The next weight, 6, is greater than 5, so 6 is not in the
knapsack. Continuing this process will show that both 2 and 3 are in the knapsack
and the total weight is brought to 0, which indicates that a solution has been found.
Were this a Merkle-Hellman knapsack encryption block, the plaintext that resulted
from a ciphertext value of 70 would be 110101.
Non-superincreasing, or normal, knapsacks are hard problems; they have no
known quick algorithm. The only known way to determine which items are in the




Figure 19.1 Encryption with knapsacks.
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19 Public-Key Algorithms
CHAPTER



knapsack is to methodically test possible solutions until you stumble on the correct
one. The fastest algorithms, taking into account the various heuristics, grow expo-
nentially with the number of possible weights in the knapsack. Add one item to the
sequence of weights, and it takes twice as long to find the solution. This is much
more difficult than a superincreasing knapsack where, if you add one more weight
to the sequence, it simply takes another operation to find the solution.
The Merkle-Hellman algorithm is based on this property. The private key is a
sequence of weights for a superincreasing knapsack problem. The public key is a
sequence of weights for a normal knapsack problem with the same solution. Merkle
and Hellman developed a technique for converting a superincreasing knapsack prob-
lem into a normal knapsack problem. They did this using modular arithmetic.
Creating the Public Key from the Prioate Key
Without going into the number theory, this is how the algorithm works: To get a
normal knapsack sequence, take a superincreasing knapsack sequence, for example
{ZWi 13,27,521, and multiply all of the values by a number n, mod m. The modulus
should be a number greater than the sum of all the numbers in the sequence: for
example, 105. The multiplier should have no factors in common with the modulus:
for example, 3 1. The normal knapsack sequence would then be
2 31 mod 105 = 62
l


3*31mod105=93
6*31mod105=81
13*31mod105=88
27 31 mod 105 = 102
l


52 * 31 mod 105 = 37
The knapsack would then be (62,93,81,88,102,37).
The superincreasing knapsack sequence is the private key. The normal knapsack
sequence is the public key.
Encryption
To encrypt a binary message, first break it up into blocks equal to the number of
items in the knapsack sequence. Then, allowing a one to indicate the item is present
and a zero to indicate that the item is absent, compute the total weights of the knap-
sacks-one for every message block.
For example, if the message were 011000110101101110 in binary, encryption
using the previous knapsack would proceed like this:
message=011000 110101 101110
011000 corresponds to 93 + 81 = 174
110101 corresponds to 62 + 93 + 88 + 37 = 280
101110correspondsto62+81+88+102=333
The ciphertext would be
174,280,333
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Decryption
A legitimate recipient of this message knows the private key: the original super-
increasing knapsack, as well as the values of n and m used to transform it into a nor-
mal knapsack. To decrypt the message, the recipient must first determine n-l such
that n(n-˜) = 1 (mod m). Multiply each of the ciphertext values by n-I mod m, and
then partition with the private knapsack to get the plaintext values.
In our example, the superincreasing knapsack is (2,3,6,13,27,52], m is equal to
105, and n is equal to 31. The ciphertext message is 174,280,333. In this case n-™ is
equal to 61, so the ciphertext values must be multiplied by 61 mod 105.
174 61 mod 105 = 9 = 3 + 6, which corresponds to 011000
l


280 61 mod 105 = 70 = 2 + 3 + 13 + 52, which corresponds to 110101
l


333 61 mod 105 = 48 = 2 + 6 + 13 + 27, which corresponds to 101110
l



The recovered plaintext is 011000 110101 101110.
Practical Implementations
With a knapsack sequence of only six items, it™s not hard to solve the problem
even if it isn™t superincreasing. Real knapsacks should contain at least 250 items.
The value for each term in the superincreasing knapsack should be somewhere
between 200 and 400 bits long, and the modulus should be somewhere between 100
to 200 bits long. Real implementations of the algorithm use random-sequence gen-
erators to produce these values.
With knapsacks like that, it™s futile to try to solve them by brute force. If a com-
puter could try a million possibilities per second, trying all possible knapsack values
would take over 1O46 years. Even a million machines working in parallel wouldn™t
solve this problem before the sun went nova.
of Knapsacks
Security
It wasn™t a million machines that broke the knapsack cryptosystem, but a pair of
cryptographers. First a single bit of plaintext was recovered [725]. Then, Shamir
showed that knapsacks can be broken in certain circumstances [ 1415,1416]. There
were other results-[ 1428,38,754,5 16,488]-but no one could break the general
Merkle-Hellman system. Finally, Shamir and Zippel [ 1418,1419,1421] found flaws
in the transformation that allowed them to reconstruct the superincreasing knap-
sack from the normal knapsack. The exact arguments are beyond the scope of this
book, but a nice summary of them can be found in [1233,1244]. At the conference
where the results were presented, the attack was demonstrated on stage using an
Apple II computer [492,494].
Knapsack Variants
Since the original Merkle-Hellman scheme was broken, many other knapsack sys-
tems have been proposed: multiple iterated knapsacks, Graham-Shamir knapsacks,
and others. These have all been analyzed and broken, generally using the same cryp-
tographic techniques, and litter the cryptographic highway [260,253,269,921,15,919,
920,922,366,254,263,255]. Good overviews of these systems and their cryptanalyses
can be found in [267,479,257,268].
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Other algorithms have been proposed that use ideas similar to those used in knap-
sack cryptosystems, but these too have been broken. The Lu-Lee cryptosystem
[990,13] was broken in [20,614,873]; a modification [507] is also insecure [1620].
Attacks on the Goodman-McAuley cryptosystem are in [646,647,267,268]. The
Pieprzyk cryptosystem [1246] can be broken by similar attacks. The Niemi cryp-
tosystem [1169], based on modular knapsacks, was broken in [345,788]. A newer
multistage knapsack [747] has not yet been broken, but I am not optimistic. Another
variant is [294].
While a variation of the knapsack algorithm is currently secure-the Chor-Rivest
knapsack [356], despite a “specialized attack” [743]-the amount of computation
required makes it far less useful than the other algorithms discussed here. A variant,
called the Powerline System, is not secure [958]. Most important, considering the
ease with which all the other variations fell, it doesn™t seem prudent to trust them.
Patents
The original Merkle-Hellman algorithm is patented in the United States [720] and
worldwide (see Table 19.1). Public Key Partners (PKP) licenses the patent, along
with other public-key cryptography patents (see Section 25.5). The U.S. patent will
expire on August 19, 1997.


19.3 RSA
Soon after Merkle™s knapsack algorithm came the first full-fledged public-key algo-
rithm, one that works for encryption and digital signatures: RSA [ 1328,1329]. Of all
the public-key algorithms proposed over the years, RSA is by far the easiest to
understand and implement. (Martin Gardner published an early description of the
algorithm in his “Mathematical Games” column in Scientific American [599].) It is


Table 19.1
Foreign Merkle-Hellman Knapsack Patents
Countrv Number Date of Issue
Belgium 871039 5 Apr 1979
Netherlands 7810063 10 Apr 1979
Great Britain 2006580 2 May 1979
Germany 2843583 10 May 1979
Sweden 7810478 14 May 1979
France 2405532 8 Jun 1979
Germany 2843583 3 Jun 1982
Germany 2857905 15 Jul 1982
Canada 1128159 20 Jul 1982
Great Britain 2006580 18 Aug 1982
Switzerland 63416114 14 Jan 1983
Italv 1099780 28 Ser,1985
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19.3 RSA


also the most popular. Named after the three inventors-Ron Rivest, Adi Shamir,
and Leonard Adleman-it has since withstood years of extensive cryptanalysis.
Although the cryptanalysis neither proved nor disproved RSA™s security, it does sug-
gest a confidence level in the algorithm.
RSA gets its security from the difficulty of factoring large numbers. The public
and private keys are functions of a pair of large (100 to 200 digits or even larger)
prime numbers. Recovering the plaintext from the public key and the ciphertext is
conjectured to be equivalent to factoring the product of the two primes.
To generate the two keys, choose two random large prime numbers, p and q. For
maximum security, choose p and q of equal length. Compute the product:
n=pq
Then randomly choose the encryption key, e, such that e and (p - l)(q - 1) are rela-
tively prime. Finally, use the extended Euclidean algorithm to compute the decryp-
tion key, d, such that
ed= 1 mod(p- l)(q- 1)
In other words,
d = e-l mod ((p - l)(q - 1))
Note that d and n are also relatively prime. The numbers e and n are the public
key; the number d is the private key. The two primes, p and q, are no longer needed.
They should be discarded, but never revealed.
To encrypt a message m, first divide it into numerical blocks smaller than n (with
binary data, choose the largest power of 2 less than n). That is, if both p and q are
loo-digit primes, then n will have just under 200 digits and each message block, m,
should be just under 200 digits long. (If you need to encrypt a fixed number of
blocks, you can pad them with a few zeros on the left to ensure that they will always
be less than n.) The encrypted message, c, will be made up of similarly sized mes-
sage blocks, ci, of about the same length. The encryption formula is simply
ci=mtmodn
To decrypt a message, take each encrypted block ci and compute
mi = Cidmod n
Since
qd = (mt)d = mled= ml& - lliq - 11 1= mimiNP - lJ(q 11 m,* 1 = mi; all
+ -=
(mod n)
the formula recovers the message. This is summarized in Table 19.2.
The message could just as easily have been encrypted with d and decrypted with
e; the choice is arbitrary. I will spare you the number theory that proves why this
works; most current texts on cryptography cover it in detail.
A short example will probably go a long way to making this clearer. If p = 47 and
q= 71, then
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Table 19.2
RSA Encryption
Public Key:
n product of two primes, p and q (p and q must remain secret)
e relatively prime to (p - l)(q - 1)
Private Key:
d e-l mod((p- l)(q- 1))
Encrypting:
c=mEmodn
Decrypting:
m=cdmodn

n=pq=3337
The encryption key, e, must have no factors in common with
(p- l)(q- 1)=46 70=3220
l



Choose e (at random) to be 79. In that case
d = 79-l mod 3220 = 1019
This number was calculated using the extended Euclidean algorithm (see Section
11.3). Publish e and n, and keep d secret. Discard p and q.
To encrypt the message
m=6882326879666683
first break it into small blocks. Three-digit blocks work nicely in this case. The mes-
sage is split into six blocks, m,, in which
ml = 688
m2 = 232
m3 = 687
m4 = 966
m5 = 668
m6 = 003
The first block is encrypted as
68879mod 3337 = 1570 = c 1
Performing the same operation on the subsequent blocks generates an encrypted
message:
c=15702756209122762423 158
Decrypting the message requires performing the same exponentiation using the
decryption key of 1019, so
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19.3 RSA


15701019
mod 3337 = 688 = ml
The rest of the message can be recovered in this manner.
RSA in Hardware
Much has been written on the subject of hardware implementations of RSA [ 13 14,
1474,1456,1316,1485,874,1222,87,1410,1409,1343,998,367,1429,523,772]. Good sur-
vey articles are [258,872]. Many different chips perform RSA encryption [1310,252,
1101,1317,874,69,737,594,1275,1563,509,1223]. A partial list of currently available
RSA chips, from [ 150,258], is listed in Table 19.3. Not all are available on the open
market.
of RSA
Speed
In hardware, RSA is about 1000 times slower than DES. The fastest VLSI hard-
ware implementation for RSA with a 51%bit modulus has a throughput of 64 kilo-
bits per second [ZSS]. There are also chips that perform 1024bit RSA encryption.
Currently chips are being planned that will approach 1 megabit per second using a
512-bit modulus; they will probably be available in 1995. Manufacturers have also
implemented RSA in smart cards; these implementations are slower.
In software, DES is about 100 times faster than RSA. These numbers may change
slightly as technology changes, but RSA will never approach the speed of symmet-
ric algorithms. Table 19.4 gives sample software speeds of RSA [918].
Software Speedups
RSA encryption goes much faster if you™re smart about choosing a value of e. The
three most common choices are 3, 17, and 65537 (216 1). (The binary representation
+
of 65537 has only two ones, so it takes only 17 multiplications to exponentiate.)
X.509 recommends 65537 [304], PEM recommends 3 [76], and PKCS #l (see Section
24.14) recommends 3 or 65537 [ 13451. There are no security problems with using


Table 19.3
Existing RSA Chips
Clock Cycles
Clock Baud Rate Bits per Number of
Per 512 Bit
Company Per 512 Bits Transistors
Speed Technology Chip
Encryption
2 micron 1024 180,000
Alpha Techn. 25 MHz 13 K .98 M
1.5 micron 298 100,000
AT&T 15MHz 19 K .4 M
2.5 micron 2.56
British Telecom 10MHz 5.1 K 1M
Gate Array 32
Business Sim. Ltd. SMHZ 3.8 K .67 M
2 micron 593 95,000
Calmos Syst. Inc. 20 MHz 28 K .36 M
1 micron 1024 100,000
CNET 25 MHz 5.3 K 2.3 M
Cryptech 14MHz 17K 120
.4M Gate Array 33,000
1.5 micron 1024 150,000
Cylink 30 MHz 6.8 K 1.2M
1.4 micron 512 160,000
GEC Marconi 25 MHz 10.2 K .67 M
1 micron 1024 400,000
Pijnenburg 25 MHz 50 K .256 M
2 micron 272 86,000
Sandia 8MHz 10 K .4 M
1 micron 512 60,000
Siemens 5MHz 8.5 K .3 M
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Table 19.4
RSA Speeds for Different Modulus Lengths
with an &bit Public Key (on a SPARC II)
1,024 bits
512 bits 768 bits
Encrypt 0.03 set 0.05 set 0.08 set
Decrypt 0.16 set 0.48 set 0.93 set
Sign 0.16 set 0.52 set 0.97 set
Verify 0.02 set 0.07 set 0.08 set



any of these three values for e [assuming you pad messages with random values-see
later section), even if a whole group of users uses the same value for e.
Private key operations can be speeded up with the Chinese remainder theorem if
you save the values of p and q, and additional values such as d mod (p - l), d mod
(q - l), and (7-l mod p [ 1283,1276]. These additional numbers can easily be calcu-
lated from the private and public keys.

of RSA
Security
The security of RSA depends wholly on the problem of factoring large numbers.
Technically, that™s a lie. It is conjectured that the security of RSA depends on the
problem of factoring large numbers. It has never been mathematically proven that
you need to factor n to calculate m from c and e. It is conceivable that an entirely
different way to cryptanalyze RSA might be discovered. However, if this new way
allows the cryptanalyst to deduce d, it could also be used as a new way to factor
large numbers. I wouldn™t worry about it too much.
It is also possible to attack RSA by guessing the value of (p - l)(q - 1). This attack
is no easier than factoring n [ 16161.
For the ultraskeptical, some RSA variants have been proved to be as difficult as
factoring (see Section 19.5). Also look at [36], which shows that recovering even cer-
tain bits of information from an RSA-encrypted ciphertext is as hard as decrypting
the entire message.
Factoring n is the most obvious means of attack. Any adversary will have the
public key, e, and the modulus, n. To find the decryption key, d, he has to factor n.
Section 11.4 discusses the current state of factoring technology. Currently, a 129-
decimal-digit modulus is at the edge of factoring technology. So, n must be larger
than that. Read Section 7.2 on public key length.
It is certainly possible for a cryptanalyst to try every possible d until he stumbles on
the correct one. This brute-force attack is even less efficient than trying to factor n.
From time to time, people claim to have found easy ways to break RSA, but to
date no such claim has held up. For example, in 1993 a draft paper by William Payne
proposed a method based on Fermat™s little theorem [1234]. Unfortunately, this
method is also slower than factoring the modulus.
There™s another worry. Most common algorithms for computing primes p and q
are probabilistic; what happens if p or q is composite? Well, first you can make the
odds of that happening as small as you want. And if it does happen, the odds are that
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19.3 RSA


encryption and decryption won™t work properly-you™ll notice right away. There are
a few numbers, called Carmichael numbers, which certain probabilistic primality
algorithms will fail to detect. These are exceedingly rare, but they are insecure [746].
Honestly, I wouldn™t worry about it.
Chosen Ciphertext Attack against RSA
Some attacks work against the implementation of RSA. These are not attacks
against the basic algorithm, but against the protocol. It™s important to realize that
it™s not enough to use RSA. Details matter.
Scenario 1: Eve, listening in on Alice™s communications, manages to collect a
ciphertext message, c, encrypted with RSA in her public key. Eve wants to be able
to read the message. Mathematically, she wants m, in which
m = cd
To recover m, she first chooses a random number, r, such that r is less than n. She
gets Alice™s public key, e. Then she computes
x=Pmodn
y=xc modn
t=r-˜modn
Ifx=Pmodn, thenr=Xdmodn.
Now, Eve gets Alice to sign y with her private key, thereby decrypting y. (Alice
has to sign the message, not the hash of the message.) Remember, Alice has never
seen y before. Alice sends Eve
u=ydmodn
Now, Eve computes
tumodn=r-˜ydmodn=r-˜Xdcdmodn=cdmodn=m
Eve now has m.
Scenario 2: Trent is a computer notary public. If Alice wants a document nota-
rized, she sends it to Trent. Trent signs it with an RSA digital signature and sends it
back. (No one-way hash functions are used here; Trent encrypts the entire message
with his private key.)
Mallory wants Trent to sign a message he otherwise wouldn™t. Maybe it has a
phony timestamp; maybe it purports to be from another person. Whatever the rea-
son, Trent would never sign it if he had a choice. Let™s call this message m™.
First, Mallory chooses an arbitrary value x and computes y = Xemod n. He can eas-
ily get e; it™s Trent™s public key and must be public to verify his signatures. Then he
computes m = ym™ mod n, and sends m to Trent to sign. Trent returns rnld mod n.
Now Mallory calculates (md mod n)x-™ mod n, which equals dd mod n and is the sig-
nature of m™.
Actually, Mallory can use several methods to accomplish these same things
[423,458,486]. The weakness they all exploit is that exponentiation preserves the
multiplicative structure of the input. That is:
[xm)d mod n = xdmd mod n
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Scenario 3: Eve wants Alice to sign m3. She generates two messages, ml and m2,
such that
m3 = m1m2 (mod n)
If Eve can get Alice to sign ml and m2, she can calculate ma:
d = (mid mod n)(m2” mod n)
m3

Moral: Never use RSA to sign a random document presented to you by a stranger.
Always use a one-way hash function first. The IS0 9796 block format prevents this
attack.
Common Modulus Attack on RSA
A possible RSA implementation gives everyone the same n, but different values
for the exponents e and d. Unfortunately, this doesn™t work. The most obvious prob-
lem is that if the same message is ever encrypted with two different exponents (both
having the same modulus), and those two exponents are relatively prime (which
they generally would be), then the plaintext can be recovered without either of the
decryption exponents [ 14571.
Let m be the plaintext message. The two encryption keys are el and e2.The com-
mon modulus is n. The two ciphertext messages are:
cl = meI mod n
c2 = meI mod n
The cryptanalyst knows n, el, e2, cl, and c2. Here™s how he recovers m.
Since el and e2 are relatively prime, the extended Euclidean algorithm can find r
and s, such that
rel + se2= 1
Assuming r is negative (either r or s has to be, so just call the negative one r), then
the extended Euclidean algorithm can be used again to calculate cl-˜. Then
(˜l-˜)-˜ * C2” = m mod n
There are two other, more subtle, attacks against this type of system. One attack
uses a probabilistic method for factoring n. The other uses a deterministic algorithm
for calculating someone™s secret key without factoring the modulus. Both attacks
are described in detail in [449].
Moral: Don™t share a common n among a group of users.
Low Encryption Exponent Attack against RSA
RSA encryption and signature verification are faster if you use a low value for e,
but that can also be insecure [704]. If you encrypt e(e + 1j/2 linearly dependent mes-
sages with different public keys having the same value of e, there is an attack
against the system. If there are fewer than that many messages, or if the messages
are unrelated, there is no problem. If the messages are identical, then e messages are
enough. The easiest solution is to pad messages with independent random values.
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19.3 RSA


This also ensures that me mod n #me. Most real-world RSA implementations-PEM
and PGP (see Sections 24.10 and 24.12), for example-do this.
Moral: Pad messages with random values before encrypting them; make sure m is
about the same size as n.
Low Decryption Exponent Attack against RSA
Another attack, this one by Michael Wiener, will recover d, when d is up to one
quarter the size of n and e is less than n [ 15961.This rarely occurs if e and d are cho-
sen at random, and cannot occur if e has a small value.
Moral: Choose a large value for d.

Lessons Learned
Judith Moore lists several restrictions on the use of RSA, based on the success of
these attacks [ 1114,1115]:

- Knowledge of one encryption/decryption pair of exponents for a given
modulus enables an attacker to factor the modulus.
- Knowledge of one encryption/decryption pair of exponents for a given
modulus enables an attacker to calculate other encryption/
decryption pairs without having to factor n.
- A common modulus should not be used in a protocol using RSA in a
communications network. (This should be obvious from the previous
two points.)
- Messages should be padded with random values to prevent attacks on
low encryption exponents.
- The decryption exponent should be large.

Remember, it is not enough to have a secure cryptographic algorithm. The entire
cryptosystem must be secure, and the cryptographic protocol must be secure. A fail-
ure in any of those three areas makes the overall system insecure.

Attack on Encrypting and Signing with RSA
It makes sense to sign a message before encrypting it (see Section 2.7), but not
everyone follows this practice. With RSA, there is an attack against protocols that
encrypt before signing [48].
Alice wants to send a message to Bob. First she encrypts it with Bob™s public key;
then she signs it with her private key. Her encrypted and signed message looks like:
(meBmod ng)dAmod nA
Here™s how Bob can claim that Alice sent him m™ and not m. Realize that since
Bob knows the factorization of nB (it™s his modulus], he can calculate discrete loga-
rithms with respect to nB. Therefore, all he has to do is to find an x such that
m™x = m mod nB
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Then, if he can publish xeg as his new public exponent and keep nB as his modu-
lus, he can claim that Alice sent him message m™ encrypted in this new exponent.
This is a particularly nasty attack in some circumstances. Note that hash func-
tions don™t solve the problem. However, forcing a fixed encryption exponent for
every user does.
Standards
RSA is a de facto standard in much of the world. The IS0 almost, but not quite,
created an RSA digital-signature standardj RSA is in an information annex to IS0
9796 [762]. The French banking community standardized on RSA [525], as have the
Australians [1498]. The United States currently has no standard for public-key
encryption, because of pressure from the NSA and patent issues. Many U.S. compa-
nies use PKCS (see Section 24.141, written by RSA Data Security, Inc. A draft ANSI
banking standard specifies RSA [61].
Patents
The RSA algorithm is patented in the United States [1330], but not in any other
country. PKP licenses the patent, along with other public-key cryptography patents
(see Section 25.5). The U.S. patent will expire on September 20, 2000.


19.4 POHLIG-HELLMAN
The Pohlig-Hellman encryption scheme [1253] is similar to RSA. It is not a sym-
metric algorithm, because different keys are used for encryption and decryption. It
is not a public-key scheme, because the keys are easily derivable from each other;
both the encryption and decryption keys must be kept secret.
Like RSA,
C=P”modn
P=Cdmodn
where
ed = 1 (mod some complicated number)
Unlike RSA, n is not defined in terms of two large primes, it must remain part of
the secret key. If someone had e and n, they could calculate d. Without knowledge
of e or d, an adversary would be forced to calculate
e = log& mod n
We have already seen that this is a hard problem.
Patents
The Pohlig-Hellman algorithm is patented in the United States (7221 and also in
Canada. PKP licenses the patent, along with other public-key cryptography patents
(see Section 25.5).
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19.5 Rabin


19.5 RABIN

Rabin™s scheme [1283,1601] gets its security from the difficulty of finding square
roots modulo a composite number. This problem is equivalent to factoring. Here is
one implementation of this scheme.
First choose two primes, p and 4, both congruent to 3 mod 4. These primes are the
private key; the product n = pq is the public key.
To encrypt a message, M (M must be less than n), simply compute
C=M2modn
Decrypting the message is just as easy, but slightly more annoying. Since the
receiver knows p and 4, he can solve the two congruences using the Chinese
remainder theorem. Compute
ml = CfJ™+ lV4mod p
m2 = (p - CfP+ “j4) mod p
m3 = Cl4 + lV4mod q
m4 = (q - C(q+ ˜li4) mod 4
Then choose an integer a = q(& mod p) and a integer b = p(p-l mod 4). The four
possible solutions are:
Ml = (amI + bm,) mod n
M2 = (urn1 + bm,) mod n
M3 = (urn2 + bm,) mod n
M4 = (am2 + bm,) mod n
One of those four results, Ml, M2, M3, or n/r,, equals M. If the message is English
text, it should be easy to choose the correct Mi. On the other hand, if the message is
a random-bit stream (say, for key generation or a digital signature), there is no way
to determine which Ml is correct. One way to solve this problem is to add a known
header to the message before encrypting.
Williams
Hugh Williams redefined Rabin™s schemes to eliminate these shortcomings [ 16011.
In his scheme, p and q are selected such that
p=3mod8
q=7mod8
and
N=pq
Also, there is a small integer, S, such that J(S,N) = -1. (J is the Jacobi symbol-see
Section 11.3). N and S are public. The secret key is k, such that
k= l/2 * (l/4 (p- 1) (q-l)+ 1)
l l
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To encrypt a message M, compute cl such that J(M,N) = (-1)“1. Then, compute M
= (SC1 M) mod N. Like Rabin™s scheme, C = M™2 mod N. And c2= 111™ mod 2. The final
l


ciphertext message is the triple:


To decrypt C, the receiver computes M“ using
Ck = +X (mod N)
The proper sign of M” is given by c2. Finally,
M = 191 * 1-l)“™ M”) mod N
l



Williams refined this scheme further in [ 1603,1604,1605]. Instead of squaring the
plaintext message, cube it. The large primes must be congruent to 1 mod 3; other-
wise the public and private keys are the same. Even better, there is only one unique
decryption for each encryption.
Both Rabin and Williams have an advantage over RSA in that they are provably as
secure as factoring. However, they are completely insecure against a chosen-
ciphertext attack. If you are going to use these schemes in instances where an
attacker can mount this attack (for example, as a digital signature algorithm where
an attacker can choose messages to be signed), be sure to use a one-way hash func-
tion before signing. Rabin suggested another way of defeating this attack: Append a
different random string to each message before hashing and signing. Unfortunately,
once you add a one-way hash function to the system it is no longer provably as
secure as factoring [628], although adding hashing cannot weaken the system in any
practical sense.
Other Rabin variants are [972,909,696,697,1439,989]. A two-dimensional variant
is in [866,889].


19.6 ELGAMAL
The ElGamal scheme [5 18,519] can be used for both digital signatures and encryp-
tion; it gets its security from the difficulty of calculating discrete logarithms in a
finite field.
To generate a key pair, first choose a prime, p, and two random numbers, g and x,
such that both g and x are less than p. Then calculate
y=g”modp
The public key is y, g, and p. Both g and p can be shared among a group of users.
The private key is x.
ElGamal Signatures
To sign a message, M, first choose a random number, k, such that k is relatively
prime to p - 1. Then compute
a=gkmodp
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19.6 ElGamal


and use the extended Euclidean algorithm to solve for b in the following equation:
M=(xa+kb)mod(p-1)
The signature is the pair: a and b. The random value, k, must be kept secret.
To verify a signature, confirm that
yaab mod p = gM mod p
Each ElGamal signature or encryption requires a new value of k, and that value
must be chosen randomly. If Eve ever recovers a k that Alice used, she can recover
Alice™s private key, x. If Eve ever gets two messages signed or encrypted using the
same k, even if she doesn™t know what it is, she can recover x.
This is summarized in Table 19.5.
For example, choose p = 11 and g = 2. Choose private key x = 8. Calculate
y=g”modp=2*mod11=3
Thepublickeyisy=3,g=2,andp=ll.
To authenticate M = 5, first choose a random number k = 9. Confirm that gcd(9,lO)
= 1. Compute
a=gkmodp=2™mod 11=6
and use the extended Euclidean algorithm to solve for b:
M=(ax+kb)mod(p-1)
5 = (8 6 + 9 b) mod 10
l l



The solution is b = 3, and the signature is the pair: a = 6 and b = 3.



Table 19.5
ElGamal Signatures
Public Key:
p prime (can be shared among a group of users)
g < p (can be shared among a group of users)
y =g”modp
Private Key:
x <P
Signing:
k choose at random, relatively prime to p - 1
a (signature) = gk mod p
b (signature) such that M = (xa + kb) mod (p - 1)
Verifying:
Accept as valid if y”ab mod p = 9” mod p
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To verify a signature, confirm that
yaab mod p = gM mod p
3663mod 11 = 25 mod 11
A variant of ElGamal for signatures is in [ 13771. Thomas Beth invented a variant
of the ElGamal scheme suitable for proofs of identity [146]. There are variants for
password authentication [312], and for key exchange [773]. And there are thousands
more (see Section 20.4).
ElGamal Encryption
A modification of ElGamal can encrypt messages. To encrypt message M, first
choose a random k, such that k is relatively prime to p - 1. Then compute
a=gkmodp
b=ykMmodp
The pair, a and b, is the ciphertext. Note that the ciphertext is twice the size of the
plaintext.
To decrypt a and b, compute
M = b/a” mod p
Since ax = gk” (mod p), and b/a” = ykM/a” = gxkM/gxk = M (mod p), this all works
(see Table 19.6). This is really the same as Diffie-Hellman key exchange (see Section
22. l), except that y is part of the key, and the encryption is multiplied by yk.
Speed
Table 19.7 gives sample software speeds of ElGamal[918].


Table 19.6
ElGamal Encryption
Public Key:
p prime (can be shared among a group of users)
g < p (can be shared among a group of users]
y =g”modp
Private Key:
x <P
Encrypting:
k choose at random, relatively prime to p - 1.
a (ciphertext) = g” mod p
b (ciphertext) = ykM mod p
Decrypting:
M (plaintext) = b/a” mod p
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19.7 McEliece


Patents
ElGamal is unpatented. But, before you go ahead and implement the algorithm,
realize that PKP feels that this algorithm is covered under the Diffie-Hellman patent
[718]. However, the Diffie-Hellman patent will expire on April 29, 1997, making
ElGamal the first public-key cryptography algorithm suitable for encryption and
digital signatures unencumbered by patents in the United States. I can hardly wait.


19.7 MCELIECE
In 1978 Robert McEliece developed a public-key cryptosystem based on algebraic
coding theory [1041]. The algorithm makes use of the existence of a class of error-
correcting codes, known as Goppa codes. His idea was to construct a Goppa code
and disguise it as a general linear code. There is a fast algorithm for decoding Goppa
codes, but the general problem of finding a code word of a given weight in a linear
binary code is NP-complete. A good description of this algorithm can be found in
[1233]; see also [1562]. Following is just a quick summary.
Let &(x,y) denote the Hamming distance between x and y. The numbers n, k, and
t are system parameters.
The private key has three parts: G™ is a k * n generator matrix for a Goppa code
that can correct t errors. P is an n * n permutation matrix. S is a k k nonsingular
l


matrix.
The public key is a k * n matrix G: G = SG™l?
Plaintext messages are strings of k bits, in the form of k-element vectors over GF(2).
To encrypt a message, choose a random n-element vector over GF(2), z, with Ham-
ming distance less than or equal to t.
c=mG+z
To decrypt the ciphertext, first compute c™ = cP™. Then, using the decoding algo-
rithm for the Goppa code, find m™ such that &(m™G, c™) is less than or equal to t.
Finally, compute m = m™S-I.
In his original paper, McEliece suggested that n = 1024, t = 50, and k = 524. These
are the minimum values required for security.


Table 19.7
ElGamal Speeds for Different
Modulus Lengths with a 160-bit
Exponent (on a SPAFK II)
512 bits 768 bits 1024 bits
Encrypt 0.33 set 0.80 set 1.09 set
Decrypt 0.24 set 0.58 set 0.77 set
Sign 0.25 set 0.47 set 0.63 set
Verify 1.37 set 5.12 set 9.30 set
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Although the algorithm was one of the first public-key algorithms, and there were
no successful cryptanalytic results against the algorithm, it has never gained wide
acceptance in the cryptographic community. The scheme is two to three orders of
magnitude faster than RSA, but has some problems. The public key is enormous: 219
bits long. The data expansion is large: The ciphertext is twice as long as the plaintext.
Some attempts at cryptanalysis of this system can be found in [8,943,1559,306].
None of these were successful in the general case, although the similarity between
the McEliece algorithm and knapsacks worried some.
In 1991, two Russian cryptographers claimed to have broken the McEliece system
with some parameters 18821.Their paper contained no evidence to substantiate their
claim, and most cryptographers discount the result. Another Russian attack, one
that cannot be used directly against the McEliece system, is in [ 1447,1448]. Exten-
sions to McEliece can be found in [424,1227,976].

Other Algorithms Based on Linear Error-Correcting Codes
The Niederreiter algorithm [ 11671is closely related to the McEliece algorithm, and
assumes that the public key is a random parity-check matrix of an error-correcting
code. The private key is an efficient decoding algorithm for this matrix.
Another algorithm, used for identification and digital signatures, is based on
syndrome decoding [ 15011; see [306] for comments. An algorithm based on error-
correcting codes [ 16211 is insecure [698,33,31,1560,32].


19.8 CURVE
ELLIPTIC CRYPTOSYSTEMS
Elliptic curves have been studied for many years and there is an enormous amount
of literature on the subject. In 1985, Neal Koblitz and V. S. Miller independently pro-
posed using them for public-key cryptosystems [867,1095]. They did not invent a
new cryptographic algorithm with elliptic curves over finite fields, but they imple-
mented existing public-key algorithms, like Diffie-Hellman, using elliptic curves.
Elliptic curves are interesting because they provide a way of constructing “ele-
ments” and “rules of combining” that produce groups. These groups have enough
familiar properties to build cryptographic algorithms, but they don™t have certain
properties that may facilitate cryptanalysis. For example, there is no good notion of
“smooth” with elliptic curves. That is, there is no set of small elements in terms of
which a random element has a good chance of being expressed by a simple algo-
rithm. Hence, index calculus discrete logarithm algorithms do not work. See [ 10951
for more details.
Elliptic curves over the finite field GF(2”) are particularly interesting. The arith-
metic processors for the underlying field are easy to construct and are relatively sim-
ple to implement for n in the range of 130 to 200. They have the potential to provide
faster public-key cryptosystems with smaller key sizes. Many public-key algo-
rithms, like Diffie-Hellman, ElGamal, and Schnorr, can be implemented in elliptic
curves over finite fields.
The mathematics here are complex and beyond the scope of this book. Those
interested in this topic are invited to read the two references previously mentioned,
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and the excellent book by Alfred Menezes [1059]. Two analogues of RSA work in
elliptic curves (890,454]. Other papers are [23,119,1062,869,152,871,892,25,895,353,
1061,26,913,914,915]. Elliptic curve cryptosystems with small key lengths are dis-
cussed in [701]. Next Computer Inc.˜s Fast Elliptic Encryption (FEE) algorithm also
uses elliptic curves [388]. FEE has the nice feature that the private key can be any
easy-to-remember string. There are proposed public-key cryptosystems using hyper-
elliptic curves [868,870,1441,1214].


19.9 LUC
Some cryptographers have developed generalizations of RSA that use various per-
mutation polynomials instead of exponentiation. A variation called Kravitz-Reed,
using irreducible binary polynomials [898], is insecure [451,589]. Winfried Miiller
and Wilfried Nbbauer use Dickson polynomials [ 1127,1128,965]. Rudolph Lid1 and
Miiller generalized this approach in [966,1126] (a variant is called the RCidi scheme),
and Nobauer looked at its security in [ 1172,1173]. (Comments on prime generation
with Lucas functions are in [969,967,968,598].) Despite all of this prior art, a group
of researchers from New Zealand managed to patent this scheme in 1993, calling it
LUC [ 1486,521,1487].
The nth Lucas number, V,(r! l), is defined as
V,(P,1)=PV,-l(P,1)-V,-z(P,11
There™s a lot more theory to Lucas numbers; I™m ignoring all of it. A good theoret-
ical treatment of Lucas sequences is in [ 1307,1308]. A particularly nice description
of the mathematics of LUC is in [ 1494,708].
In any case, to generate a public-key/private-key key pair, first choose two large
primes, p and 4. Calculate n, the product of p and 4. The encryption key, e, is a ran-
dom number that is relatively prime to p - 1, 4 - 1, p + 1, and 4 + 1.
There are four possible decryption keys,
d = e-l mod (lcm((p + l), (4 + 1)))
d = e-l mod (lcm((p + l), (q - 1)))
d = e-l mod (lcm((p - l), (4 + 1)))
d = e-l mod (lcm((p - l), (4 - 1)))
where lcm is the least common multiple.
The public key is d and n; the private key is e and n. Discard p and 4.
To encrypt a message, P (P must be less than n), calculate
C = V,(r! 1) (mod n)
And to decrypt:
P = Vd(P,1) (mod n), with the proper d
At best, LUC is no more secure than RSA. And recent, still-unpublished results
show how to break LUC in at least some implementations. I just don™t trust it.
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19.10 FINITE AUTOMATON PUBLIC-KEY CRYPTOSYSTEMS
Chinese cryptographer Tao Renji has developed a public-key algorithm based on
finite automata [ 1301,1302,1303,1300,1304,666]. Just as it is hard to factor the prod-
uct of two large primes, it is also hard to factor the composition of two finite
automata. This is especially so if one or both of them is nonlinear.
Much of this research took place in China in the 1980s and was published in Chi-
nese. Renji is starting to write in English. His main result was that certain nonlin-
ear automata (the quasilinear automata) possess weak inverses if, and only if, they
have a certain echelon matrix structure. This property disappears if they are com-
posed with another automaton (even a linear one). In the public-key algorithm, the
secret key is an invertible quasilinear automaton and a linear automaton, and the
corresponding public key can be derived by multiplying them out term by term.
Data is encrypted by passing it through the public automaton, and decrypted by
passing it through the inverses of its components (in some cases provided they have
been set to a suitable initial state). This scheme works for both encryption and dig-
ital signatures.
The performance of such systems can be summed up by saying that like McEliece™s
system, they run much faster than RSA, but require longer keys. The keylength
thought to give similar security to 512-bit RSA is 2792 bits, and to 1024-bit RSA is
4152 bits. For the former case, the system encrypts data at 20,869 bytes/set and
decrypts data at 17,117 bytes/set, running on a 33 MHz 80486.
Renji has published three algorithms. The first is FAPKCO. This is a weak system
which uses linear components, and is primarily illustrative. Two serious systems,
FAPKC 1 and FAPKC2, use one linear and one nonlinear component each. The latter
is more complex, and was developed in order to support identity-based operation.
As for their strength, quite a lot of work has been done on them in China (where
there are now over 30 institutes publishing cryptography and security papers). One
can see from the considerable Chinese language literature that the problem has been
studied.
One possible attraction of FAPKCl and FAPKC2 is that they are not encumbered
by any U.S. patents. Thus, once the Diffie-Hellman patent expires in 1997, they will
unquestionably be in the public domain.