Voting protocol

Secure and Fault-Tolerant Voting in Distributed Systems

Ben Hardekopf, Kevin Kwiat, Shambhu Upadhyaya
¡
Air Force Research Laboratory
AFRL/IFGA
525 Brooks Rd.
Rome, NY 13441-4505
¢ hardekob, kwiatk£ @rl.af.mil; shambhu@cse.buffalo.edu
Abstract – Concerns about both security and fault-tolerance
have had an important impact on the design and use of dis-
tributed information systems in the past. As such systems
become more prevalent, as well as more pervasive, these con-
cerns will become even more immediately relevant. From
national defense, to commercial interests, to private citizens,
distributed systems are making an ever greater impact on our
lives.
We will focus here on integrating security and fault-tolerance
into one, general-purpose protocol for secure distributed vot-
ing. Distributed voting is a well-known fault-tolerance tech-
nique [4]. For the most part, however, security had not been
a concern in systems that used voting. More recently, several
protocols have been proposed to shore up this lack. These
protocols, however, have limitations which make them par-
ticularly unsuitable for many aerospace applications, because
those applications require very flexible voting schemes (e.g.,
voting among real-world sensor data).
We present a new, more general voting protocol that reduces
the vulnerability of the voting process to both attacks and
faults. The algorithm is contrasted with the traditional 2-
phase commit protocols typically used in distributed voting
and with other proposed secure voting schemes. Our algo-
rithm is applicable to exact and inexact voting in networks
where atomic broadcast and predetermined message delays
are present, such as local area networks. For wide area net-
works without these properties, we describe yet another ap-
proach that satisfies our goals of obtaining security and fault
tolerance for a broad range of aerospace information systems.
¤ 0-7803-6599-2/01/$10.00 c
¥ 2001 IEEE¦ Author affiliated with SUNY Buffalo; work performed while under the
Air Force Research Laboratory/Information Directorate’s 2000 Summer Fac-
ulty Research Program.
TABLE OF CONTENTS
1. INTRODUCTION
2. PAST WORK
3. THE ALGORITHM
4. DISCUSSION
5. WIDE-AREA NETWORKS
6. CONCLUSION
1. INTRODUCTION
As evidenced by recent news stories, the aerospace realm is
not immune from the concerns of security as well as fault-
tolerance. Although a recent hacking incident at NASA was
deemed not to have endangered a shuttle mission ([1]), at
the very least, it calls attention to the potential threats posed
to the aerospace community. Networking on-board systems
with those that are ground-based raises the concern of open-
ing avenues for highly detrimental attacks. To be truly com-
prehensive, dependable aerospace information systems must
tolerate faults that manifest themselves as a result of random
phenomena or deliberate interference.
Historically, aerospace missions have been among the first to
use fault tolerance. Early visionaries [2, 3] of the Apollo Pro-
gram foresaw using redundancy to combat failure: by having
a second spacecraft accompany a crew to the moon’s surface,
their return would be ensured should their primary landing
vehicle be damaged. Redundancy at this level of granular-
ity never came about because durability improvements in the
eventual design of the Lunar Module reduced the risk to the
crew of having only one of them on the moon.
Leaping to the present we see that special design techniques
have been required for the computers used in aerospace

missions. The Self-Testing And Repairing (STAR) Com-
puter (1971), the Fault-Tolerant Multiprocessor (1975), the
Fault-Tolerant Spaceborne Computer (FTSC) (1976), and the
Multi-Microprocessor Flight Control System (1981) are prac-
tical computer systems that perform critical mission functions
and have been specifically designed to ensure mission success
[4]. Redundancy of selected computer components within
these designs plays an important role in reducing the risk as-
sociated with relying upon any single component to operate
flawlessly. In one of the most highly visible applications of
fault-tolerant computing, the Space Shuttle makes use of re-
dundancy at the level of general-purpose computers to ensure
that flight-critical operations such as ascent, reentry, and land-
ing are performed in spite of the failure of any one computer.
Distributed computer systems, as an automatic consequence
of their architecture, can be configured for concurrent oper-
ation in addition to offering resilience against hardware fail-
ure [6]. This attractive dual-property was also observed by
the designers of the Space Shuttle computer in their vision
of on-board systems for advanced Space Shuttles and Space
Stations [5].
Distributed systems are important not only to on-board sys-
tems, but also to ground based systems. Tracking sys-
tems placed around the globe are inherently physically dis-
tributed; yet they have to be linked to command centers. Also,
aerospace missions may concern multiple ground-based cen-
ters that must coordinate their separate activities through
communication networks. Connectivity to this degree cre-
ates distributed systems of distributed systems. Dependency
among these systems’ components requires fault-tolerant de-
sign to combat the likelihood of mission failure due to system
component failure.
The focus of this paper is on distributed voting, a well-known
fault-tolerance technique in which multiple voters indepen-
dently compute their results and vote to determine a major-
ity; the majority result is then committed (sent to the user).
Much work has been done in this area, as described in the
next section on past work. Our contribution is a radical ap-
proach, in which the normal order of events is reversed –
any arbitrary voter can initiate a committal, but the result is
buffered long enough for the other voters to check the result
and vote on whether they should recommit a new result if the
first was incorrect. This process is repeated until a majority
of voters agree, and the final result is sent to the user. The se-
curity implications of this simple change are profound. The
next section reviews the basic concepts of distributed voting,
and presents several algorithms that have been proposed in
the past. Then our new proposal, the timed-buffer distributed
voting algorithm (TB-DVA) is described and analyzed in sec-
tions three and four. Section five presents some preliminary
work on another algorithm which operates under looser re-
strictions than the TB-DVA, and hence can be even more gen-
erally applied.
2. PAST WORK
Replication and majority voting are the conventional methods
for achieving fault tolerance in distributed systems. Decen-
tralized voting, in which the replicated voters independently
determine the majority rather than relying on a central server
to tally the results, has become the strategy of choice, and has
had a number of incarnations [7, 8, 9]. Most of these systems
have used the 2-phase commit protocol in order to implement
the voting scheme. In this protocol, the replicated voters first
exchange their votes and independently determine the major-
ity result. Once a final result has been calculated, one of the
voters is arbitrarily chosen to commit that result (i.e. to pass
the result on to the user). This method is widely advocated
in designing fault-tolerant open distributed systems [10]. The
problem with this type of protocol lies in the committal phase.
If the voter chosen to commit the result fails right before or
during the committal, the user will receive a bad result. The
probability of this happening is slight, and usually falls within
acceptable risk parameters. However, if security as well as
fault-tolerance is to be taken into account, then the problem
is greatly exacerbated. If a hostile attacker has taken control
of the committing voter, then the attacker can control what
results the user sees, regardless of the other voters’ results.
There have been several protocols proposed that attempt to
overcome this problem. For example, the algorithm presented
in [11] works as follows (in a very simplified presentation):
1. A client sends a request to one of the voters.
2. The voter multi-casts the request to the other voters.
3. The voters execute the request and send a reply to the
client.
4. The client waits for §©¨ replies from different voters
with the same result, where § is the number of faults to
be tolerated; this is the final result.
While this strategy obviously is not subject to the same prob-
lem as the 2-phase commit protocol, since in essence all the
voters commit a result, it does require substantial computa-
tion on the part of the client, which must collect and compare
all the replies until §¨ have been collected that carry the
same result. As a result, this system does not scale very well.
Another protocol that attempts to alleviate this problem is de-
scribed in [12]. It makes use of a (k,n)-threshold signature

scheme. Informally, this describes a scheme wherein a public
key is generated, along with shares of the corresponding
private key, each of which can be used to produce a partial re-
sult on a signed message . Any of these partial results can
then be used to reconstruct the whole of . In this particular
protocol, is the number of voters, and is set as one more
than the number of tolerated faults. Each voter signs its result
with its particular share of the private key and broadcasts it to
the other voters. The voter then sorts through the broadcast
messages for partial results which agree with its own result
and can be combined into the whole message , where
would be the signed final result. The voter then sends to
the client, which accepts the first such valid sent. Again,
this protocol is not subject to the error inherent in the 2-phase
commit protocol, and it is also not computationally expensive
for the client. However, it achieves this by shifting the com-
putational burden to the voters. As a result, this system also
does not scale very well.
Other protocols have also been proposed, each with its own
advantages and disadvantages [13, 14, 15, 16]. However, all
of the above schemes for securing the distributed voting pro-
cess make the common assumption which underlies the idea
of state-machine replication – two different voters, starting in
the same state and following the same instructions, will in-
evitably arrive at the same result. While there are many cases
when this assumption holds, there are also times when it does
not. This is true in the case of so-called inexact voting [17].
In inexact voting, two results do not have to be bit-wise identi-
cal in order to be considered equal, as long as they fall within
some pre-defined range of tolerance. This situation often
arises when data is gathered from sensors interacting with the
real world – it is extremely unlikely that two different sensors
will collect exactly the same data, even if they are arbitrar-
ily close to one another and sampling the same phenomena;
therefore some analysis needs to be done to determine if the
sensors’ data is effectively equal, even if not identical.
In such situations the schemes described above will encounter
problems, because of the common assumption they all make
that the replicated voters’ data will be identical. The sec-
ond algorithm described above, which uses a (k,n)-threshold
scheme, cannot be used for inexact voting – in order for the
partial results to be combined together into a whole result for
the client, the partial results must be identical.
While some of the algorithms could be modified to handle
inexact voting, the performance cost incurred through multi-
ple inexact comparisons would be prohibitive. For example,
the first algorithm described in this section, in which all vot-
ers send their results to the client, would force the client to
make multiple inexact comparisons in order to determine the
majority. Since inexact comparisons can be very complex op-
erations, this places an unacceptable burden on the client.
3. THE ALGORITHM
Assumptions
The proposed algorithm has two sets of participants. One is
the set of voters, which can be arbitrarily large but must have
at least three elements. These voters are completely indepen-
dent; the only exchange of information that takes place be-
tween them is communicating the voters’ individual results.
The other set contains the user and an interface module. The
interface module buffers the user from the voters (see Fig-
ure 1). The interface module consists, in its abstract form,
of a simple memory buffer and timer. A task is sent from
the user, through the interface module, to the voters. At the
termination of the algorithm, the interface module passes the
final result back to the user.
The environment for the algorithm is a network with an
atomic broadcast capability and bounded message delay (e.g.,
a local area network). It is assumed that a fair-use policy
is enforced, so that no host can indefinitely appropriate the
broadcast medium [18]. It is also assumed that no voter will
commit an answer until all voters are ready – this can be
easily enforced by setting an application dependent threshold
beyond which all functional voters should have their results
ready; any commits attempted before this threshold is reached
are considered automatically invalid. Each voter can commit
only once – this is enforced at the interface module, which ig-
nores commits from a voter which has previously committed.
The most important assumption made is that a majority of
the participating voters are fault-free and follow the protocol
faithfully. No assumptions are made about the remaining vot-
ers – they can refuse to participate, send arbitrary messages,
commit incorrect results, etc.; they are not bound in any way.
Description
Each of the (correct) voters will follow the steps below:
1. If no other voter has committed an answer to the inter-
face module yet, the voter does so with its own vote; it
then skips the remaining steps.
2. In the case that another voter has committed, the voter
compares the committed value from the other voter with
its own vote.
3. If the results agree, the voter does nothing; otherwise it
broadcasts its dissenting vote to all the other voters.

Interface Module
User
!#%$')( !#%$''0 !1%$''2 !#%$'43
Figure 1: System Architecture
4. Once all voters have had a chance to compare their votes
with the committed value (this interval would be deter-
mined by a timer), the voter analyzes all the dissenting
votes to determine if a majority dissenting vote exists.
5. If no majority exists, then the voter does nothing.
6. If a new majority exists (or if another, perhaps faulty,
voter commits a new result), then the voter returns to
step 1.
The interface module will follow these steps:
1. Once a commit is received, the result is stored in the
buffer and the timer is started. The timer is set to allow
time for all the voters to check the committed value and
dissent if necessary.
2. If a new commit is received before the timer runs out,
the new result is written over the old one in the buffer,
and the timer is restarted.
3. If no commit occurs before the timer runs out, then the
interface module sends the result in its buffer to the user,
and the algorithm is terminated.
4. DISCUSSION
Authentication
For the correct execution of the voting algorithm it is neces-
sary that the commits sent to the timed-buffer module from
the various voters be authenticated. Any known sophisticated
authentication techniques can be used to enforce secure com-
munication, but it should be done without increasing the com-
plexity of the buffer module. For illustrative purposes we
describe a simple authentication technique that doesn’t em-
ploy standard cryptographic methods such as public key en-
cryption [19]. The technique described here is called SKEY
authentication [19], which is simple to implement but is ca-
pable of strong authentication with minimal communication
between the voters and the interface module. This approach
allows the implementation of our secure and fault tolerant
voting on the existing platforms without any modiﬁcations
to the underlying protocols.
The SKEY authentication is based on a one-way function.
The voter and the host on which the buffer module is built
ﬁrst agree on a common random number 5 prior to the start
of the voting algorithm. A set of numbers 687@9A6)BC9%DEDFDE9G6IH is
generated at a given voter as well as the host by applying the
one-way function § on 5 as 6 7QP §SRT5VUW9G6 BXP §SRT§SRT5VUGU , and
so on. The host also calculates and stores 6 H`Ya7 . The voter
sends its commit by appending 6 H to its vote. The host will
calculate §SRb6 H U and compare it with 6 H`Ya7 . If these numbers
match, the communication is treated as authentic. The voter
will delete 6 H and use 6 H'cd7 next time when it has to commit
to the timed-buffer module.
Since the SKEY method requires only an occasional ex-
change of a random number between the voters and the host
computer in which the timed-buffer resides, a reasonable
level of security can be maintained on the exchange of votes.
The Interface Module
The function of the interface module is to record a commit
from a voter, set up a timer and wait until the timeout expires
and deliver the correct result to the user. It is possible that the

timer may be reset several times before passing the final result
to the user. In addition, the interface module should have the
capability to authenticate voters prior to their committals so
that it can track the voters to ensure that a voter can commit
only once in a given voting cycle. In order to reduce the like-
lihood of attacks on the interface, it should be isolated from
the rest of the voter complex and be built to have minimal
interaction with the outside world.
Depending upon the level of voting, the design of the inter-
face module may vary. Voting may proceed at either hard-
ware or software levels. It essentially depends on the volume
of data, complexity of computation, approximation and con-
text dependency of the voting algorithms. If low level, high
frequency voting is to be done, a hardware implementation is
preferred; if high level voting with low frequency is desired,
a software implementation of the interface module may be
suitable. This is because the voting is generally much more
complex at higher levels of abstraction. We assume low level,
high frequency voting in hardware and discuss a hardware ar-
chitecture for the interface module below.
Since only one copy of the vote needs to be buffered before
giving it to the user, the amount of memory required is small.
The actual size depends on the data that is voted upon. The
tracking of voters can be implemented using a flag register.
One bit flag per voter is sufficient. The flag will be set as soon
as a commit is received from a voter and will be reset after the
expiry of the timer. If multiple commits are received from the
same voter during the flag set state, they will be ignored.
A small amount of additional memory must be built into the
interface module to support the SKEY authentication of com-
munication between the voters and the timed-buffer as de-
scribed before. This memory is needed to store a sequence
of numbers for each of the voters as required by the SKEY
method of authentication. Control logic must be designed
into the interface module to step down the sequence each time
a commit is received from a voter. Re-initialization of the se-
quence for a specific voter is necessary when the sequence
reduces to zero, over time. This can be done by requesting
the host computer to receive a new random number from the
voter and computing a new sequence. Another capability that
needs to be built into the interface module is the synchroniza-
tion of result delivery with the expiry of the timer.
Though the interface module may be viewed as a single point
of failure, it is far less vulnerable to failure than a voter would
be due to the decreased level of complexity compared to the
voter/processor module. The interface module has no require-
ment to run any algorithm (code). It is isolated from the voter
complex and is designed to have minimal hardware and min-
imal interaction with the outside world. Thus, it is less vul-
nerable to attacks as well.
Correctness
The algorithm described here has been formally specified in
Lamport’s Temporal Logic of Action [20], and verified to be
correct. The proof of this result, however, is beyond the scope
of this paper.
Performance
Besides the security and fault-tolerance attributes of the al-
gorithm, another important characteristic is its performance.
A detailed analysis was performed in [21], and showed that
this algorithm had definite performance advantages. To sum-
marize the conclusions of that paper, it was determined that
this algorithm had an average efRghU performance in relation
to the number of voters used – i.e., the algorithm scales ex-
tremely well to systems with large numbers of voters. This
result is especially important given that the security and fault-
tolerance (as opposed to performance) of a system using this
algorithm rises linearly with the number of voters in the sys-
tem.
Intrusion Tolerance
Another benefit of this algorithm that we have yet to fully ex-
plore is its applicability to the problem of intrusion tolerance.
Any voter that commits an incorrect value can be partitioned
from the network and flagged for review by a higher author-
ity (either automated or human) as a possible security breach.
Assuming that all voters are denied access to covert channels
(a strong assumption), we can also have each voter monitor
all other voters, and in a similar manner flag any voter that is
releasing confidential information. We will be further explor-
ing these possibilities and what benefit they can bring to the
security of the system as a whole.
5. WIDE-AREA NETWORKS
This section discusses some preliminary work on secure inex-
act voting in a wide-area network, where the assumptions of
atomic broadcast and bounded message delay are not practi-
cal. We take advantage of Lamport’s results described in [22],
where he concludes that Byzantine fault-tolerance can be
much simplified through the use of digital signatures. Again,
the unique aspect of this algorithm, just as for the previous,
is the way it juxtaposes the requirements for security, fault-

tolerance, and performance in inexact voting.
Motivation
Distributing data and computation over a wide area network is
becoming a standard practice. Critical databases have already
been replicated and dispersed to various geographical sites to
increase their longevity [23, 24]. Redundant computations
are also distributed in order to combat localized network fail-
ures and attacks, increasing both security and fault-tolerance.
As a consequence, redundant computations on replicated data
at remote locations must somehow coordinate their results in
order to present a majority result to the user. One example
of this requirement is gathering data from distributed sensors
with overlapping areas of coverage. Determining a major-
ity result from these sensors produces the lowest probabil-
ity of error for the widest range of observation probabilities
[25]. Data need not be identical. It may even be made dif-
ferent deliberately: data diversity [26] is a software fault tol-
erance strategy where a related set of points in a program’s
data space are obtained, executed upon using the same soft-
ware, and then a decision algorithm (i.e., voter) determines
the resulting output.
Centralized voting (having a distinguished coordinator which
collects the votes from all voters and then determines the ma-
jority) is a simple solution to the problem of resolving the
output of redundant voters in a wide area network. However,
as networking becomes more ubiquitous the advantages of
distributed (i.e. decentralized) voting become clear.
Use of a centralized coordinator, which may be quite dis-
tant from the participating voters, could consume much more
bandwidth than distributed voting, in which the voters need
only communicate among themselves. Transmitting results
from the voters to the coordinator may involve many network
hops and accrue more overall delay than having the voters
communicate among themselves. Designating a node that is
close to the redundant voters to act as a ’delegate’ coordinator
may not be possible because it entails placing complete trust
in that delegate and assuring that a dependable communica-
tions link exists between it and the result’s final destination.
Another problem with centralized voting is the possibility
of link failure which may partition the network, rendering
communication between either the voters and the coordinator
or the coordinator and the user impossible. In a distributed
scheme, as long as a majority of the voters can communicate
a final result can be calculated; and as long as the user can
communicate with any of the participating voters it can ob-
tain that result.
The fact that the coordinator is receiving messages from each
and every voter makes network congestion in its vicinity
likely, especially if it is responsible for many redundant tasks
carried out at the same time (and hence is receiving messages
from many voters at once). Decentralized voting distributes
the message traffic attendant on each task and thus tends to
confine it to the participating voters.
Decentralized voting also allows the necessary computation
for determining the majority to be distributed and calculated
in parallel among the voters. Insufficient computing capac-
ity of the coordinator can restrict the usefulness of central-
ized voting. Research has been done in the area of software
agents that perform centralized voting [27], but no consider-
ation has been given to agents that may not be able to com-
pute the majority, but only apply it. Such “bounded rational
agents” have limited decision capabilities due to restrictions
placed upon them regarding the computational resources they
can consume [28]. This may be a problem when the task
of comparing two votes involves complex calculations, such
as when the votes may be somewhat different, yet still be in
agreement. Determining the majority of correct-yet-different
results calls for “inexact voting” that, being potentially far
more complex than a mere bit-wise comparison of results
[29], can readily exceed an agent’s limited decision making
power. Requiring the coordinator to correctly decide among
results that can differ but still be correct is understandable
when one considers, e.g., the tolerances of sensor readings.
Being unable to compute a majority result, an agent that ob-
tains the result from elsewhere could nonetheless use it to, for
example, manipulate an actuator through a microcontroller.
A final consideration when using centralized voting is the
possibility of an adversary observing the network. Such an
adversary could, using network traffic analysis, easily deter-
mine the importance of the coordinator from the sheer num-
ber of messages it was receiving. Being distinguished in this
manner makes the coordinator a tempting target for attack.
Once the coordinator has been compromised, the attacker has
complete control over the results seen by the user. In decen-
tralized voting, no voter is more important than any other.
Done correctly, an attacker would have to compromise a ma-
jority of the voters before being able to control the results
seen by the user, greatly increasing the cost of any successful
attack.
Issues concerning network congestion, the inability to des-
ignate alternate trustworthy coordinators, link failures, the
potentially complex computation for determining a majority,
and security all motivate the use of distributed voting in a
wide area network.
Assumptions
While the underlying wide-area network itself may be un-

reliable, we assume that this algorithm operates on top of a
reliable transport protocol, guaranteeing eventual delivery of
messages (although the messages are not necessarily deliv-
ered in the order they were sent). On top of this layer is an-
other layer which guarantees eventual delivery of valid mes-
sages - messages which have been digitally signed and cor-
rectly verified as described in the next paragraph. Messages
which cannot be verified are discarded. We assume the pres-
ence of a public-key infrastructure [19], in which each voter
has a private key and each voter knows (or can securely ob-
tain) the public key of every other voter. Each voter knows
a priori who the other voters are. We further assume that a
majority of the voters are fault-free and will correctly follow
the protocol (i.e., they are trustworthy). As before, no as-
sumptions are made about the remaining voters. There is no
interface module in this system – just the voters and the client.
The ultimate goal of the algorithm is to have each trustworthy
voter agree with every other trustworthy voter on one final re-
sult, and to have proof that its result is that which was agreed
on.
Two different functions are employed in the algorithm: one-
way hashes and digital signatures. A one-way hash is a func-
tion that maps an argument to a unique fixed-width value in
such a way that it is impossible to recover the original ar-
gument from that value. A digital signature can be accom-
plished in several ways; one mechanism is encrypting a mes-
sage (or the hash of a message) with a private key. The sig-
nature can be verified by decrypting the signature with the
corresponding public key. This provides a secure method of
authentication. All signatures include a timestamp to guard
against replay attacks.
Description
Each (correct) voter will follow the steps below:
1. Compute a result.
2. Compute the hash of the result and save that value.
3. Sign the result and send it to all the other voters.
4. For all the signed results received from the other voters:
(a) Make sure that this result isn’t a repeat (i.e., there
is only one result per voter).
(b) Verify the signature to make sure it is a valid result.
(c) If the result agrees with this voter’s result (using in-
exact comparison if necessary), then hash the other
voter’s result, sign the hash, and send it back to the
other voter (this signed hash is called an endorse-
ment).
5. For all endorsements received from the other voters:
(a) Make sure the endorsement isn’t a repeat (i.e., only
one endorsement per voter).
(b) Verify the signature and compare the hash value to
the value saved in step 2 in order to make sure it is
a valid endorsement.
6. Once a majority of endorsements has been received, the
algorithm is terminated.
The voters end up with a majority of endorsements for their
result, and once a majority vote has been determined the vot-
ers can, if necessary, transmit the result to any interested host
along with the relevant endorsements. The host can accept
the first such result accompanied by a majority of endorse-
ments which are all verified correctly, knowing that that vote
is the result agreed to by a majority of the voters. We are
guaranteed that a majority of endorsements will be received
by correct voters because of the assumption that a majority of
the voters will operate correctly.
Discussion
The goal of the algorithm, as stated earlier, is to enable vot-
ers to agree on a common result and provide proof that their
result is the one that was agreed on. It must do this in an
environment where all messages must pass through unknown
(and possible untrustworthy) intermediary nodes, and where
all of the voters are not themselves necessarily trustworthy.
The mechanism that makes this possible is the public-key dig-
ital signature. With this, voters are able to determine the orig-
inator of a message and verify that no-one tampered with the
message before it was received. This means that the interme-
diary nodes cannot influence any of the voters – they can only
relay messages (note that because of the stated assumption of
a reliable transport protocol, intermediate nodes cannot in-
definitely delay messages either). It also means that no voter
can masquerade as another voter, nor can any voter fake an
endorsement from any other voter.
In the second round of the inexact voting algorithm, signing
the hash rather than the result itself is a convenience. The re-
sult may be of any size from a simple number to a multi-field
record depending on the application, while the hash would
always be a constant size (e.g. 160 bits). If the result itself
were going to be signed as proof of correctness, then there
would be one of two options. One would be that the voters
could exchange signed votes, in which case each voter would
have to store multiple copies of the same vote, each signed by
a different voter. In order to prove to a host that the result was
correct, a voter would have to transmit each of the votes to the
host, which would in turn have to verify and compare them
all. The other option would be that the voters could each in
turn sign a vote, so that each vote would be signed multiple

Advantages Disadvantages
Centralized simple to implement single point-of-failure; rigid archi-
tecture
Distributed no single point-of-failure; flexible
architecture
complex to implement; reliance on
committing voter
Table 1: Comparison Chart for Centralized and Distributed Voting.
times. This would necessitate that the vote from each voter
be sent to a majority of the other voters, greatly expanding
the number of messages necessary. A side-benefit of using
the hash is that the intermediary nodes cannot determine the
value of the various votes they are relaying, since there is no
way to de-hash a one-way hash (hence the name). Of course,
an untrustworthy voter can relay its own result to anyone it
wishes, so this does not provide absolute confidentiality.
The requirement for timestamps for each signature is there
in order to guard against resend attacks. An attacker could
record the messages sent in a previous run of the algorithm
and resend them to the voters in a subsequent run. If there
was no way of determining that these were old messages, the
voters could be fooled into accepting them as valid votes. But
since the hashes of these votes would not match the hashes
of the voters’ results, the votes would be discarded and the
voters wouldn’t be able to agree on a majority result – even
though a majority of them may be functioning correctly.
Performance
Performance of the algorithm can be measured by the com-
plexity of the operations required of each voter and the num-
ber of messages required to be sent over the network. In the
following analysis, is the number of voters.
The first step for each voter is to calculate its result and sign
and hash that result. Since each voter does this only once,
and in parallel, this can be taken as a constant. Each voter
will then receive one signed vote from every other voter. For
each signed vote the voter must verify it and compare it with
its own result. Since this is inexact voting, this comparison
may be computationally expensive. If the vote agrees with
the voter’s result, the voter hashes and signs the vote (a trivial
operation relative to the comparison). Each voter will then re-
ceive a maximum of one endorsement from every other voter,
which they will have to verify and compare with the hash of
their own result. The complexity for each voter is therefore
efRbiU . Every voter sends one signed vote to every other voter,
resulting in pRbrqs@U messages. Each voter then sends at a
maximum one endorsement to every other voter, causing an-
other pRbtquhU messages, for a total of vwpRbtquhU messages.
Therefore the complexity of the algorithm with regards to the
number of messages is efRb B U .
6. CONCLUSION
Faults present risk to the success of an aerospace mission, so
they will continually be a concern of the fault tolerance com-
munity. We have taken up the issue of security in conjunc-
tion with fault tolerance. This motivated us to devise new
approaches to distributed voting. Within a LAN (and some
cases a WAN) we replaced the ubiquitous 2-phase commit
protocol with one that is light-weight and improves both per-
formance and security without losing any of the traditional
fault coverage. Accompanying this algorithm is one that
we proposed for resolving correct-but-possibly-not-identical
votes within a WAN. Both of these algorithms are used to
uniquely enhance the integrity of distributed information sys-
tems – protecting them from faults and hostile attacks. Ta-
ble 1 contrasts the advantages and disadvantages of central-
ized and distributed voting. The contribution of the tech-
niques described in this paper is to remove some of the dis-
advantages of distributed voting evident in the third quadrant.
Applying these algorithms to those distributed systems used
for aerospace missions can significantly contribute to the like-
lihood that the mission will succeed.
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Ben Hardekopf is a lieutenant in the United States Air Force,
stationed at the Air Force Research Laboratory. He received
the BSE in Electrical Engineering from Duke University and
is currently working towards the MS in Computer Science
from the State University of New York at Utica/Rome.
Dr. Kevin A. Kwiat has been with the U.S. Air Force Re-
search Laboratory for over 17 years. He is an adjunct profes-
sor of Computer Science at the State University of New York
at Utica/Rome, and an adjunct professor of Mathematics at
Utica College of Syracuse University. He received the BS in
Computer Science, the BA in Mathematics, the MS in Com-
puter Engineering, and the Ph.D. in Computer Engineering,
all from Syracuse University. He holds 1 patent.
Shambhu Upadhyaya received his Ph.D. in Electrical and
Computer Engineering from the University of Newcastle,
Australia in 1987. He is currently an Associate Professor of
Computer Science and Engineering at the State University of
New York at Buffalo. His research interests are fault-tolerant
computing, distributed systems, and security.

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