No Free Lunch Theorems for Optimization

No Free Lunch Theorems for Optimization
No Free Lunch Theorems for Optimization
David H. Wolpert
IBM Almaden Research Center
N5Na/D3
650 Harry Road
San Jose, CA 95120-6099
William G. Macready
Santa Fe Institute
1399 Hyde Park Road
Santa Fe, NM, 87501
December 31, 1996
Abstract
A framework is developed to explore the connection between e ective optimization
algorithms and the problems they are solving. A number of \no free lunch" (NFL)
theorems are presented that establish that for any algorithm, any elevated performance
over one class of problems is exactly paid for in performance over another class. These
theorems result in a geometric interpretation of what it means for an algorithm to
be well suited to an optimization problem. Applications of the NFL theorems to
information theoretic aspects of optimization and benchmark measures of performance
are also presented. Other issues addressed are time-varying optimization problems
and a priori \head-to-head" minimax distinctions between optimization algorithms,
distinctions that can obtain despite the NFL theorems' enforcing of a type of uniformity
over all algorithms.
1 Introduction
The past few decades have seen increased interest in general-purpose \black-box" optimization algorithms that exploit little if any knowledge concerning the optimization problem on
which they are run. In large part these algorithms have drawn inspiration from optimization
processes that occur in nature. In particular, the two most popular black-box optimization
strategies, evolutionary algorithms FOW66, Hol93] and simulated annealing KGV83] mimic
processes in natural selection and statistical mechanics respectively.
1
In light of this interest in general-purpose optimization algorithms, it has become important to understand the relationship between how well an algorithm a performs and the
optimization problem f on which it is run. In this paper we present a formal analysis that
contributes towards such an understanding by addressing questions like the following: Given
the plethora of black-box optimization algorithms and of optimization problems, how can we
best match algorithms to problems (i.e., how best can we relax the black-box nature of the
algorithms and have them exploit some knowledge concerning the optimization problem)? In
particular, while serious optimization practitioners almost always perform such matching, it
is usually on an ad hoc basis how can such matching be formally analyzed? More generally,
what is the underlying mathematical \skeleton" of optimization theory before the \esh" of
the probability distributions of a particular context and set of optimization problems are imposed? What can information theory and Bayesian analysis contribute to an understanding
of these issues? How a priori generalizable are the performance results of a certain algorithm
on a certain class of problems to its performance on other classes of problems? How should
we even measure such generalization how should we assess the performance of algorithms
on problems so that we may programmatically compare those algorithms?
Broadly speaking, we take two approaches to these questions. First, we investigate what
a priori restrictions there are on the pattern of performance of one or more algorithms as one
runs over the set of all optimization problems. Our second approach is to instead focus on
a particular problem and consider the eects of running over all algorithms. In the current
paper we present results from both types of analyses but concentrate largely on the rst
approach. The reader is referred to the companion paper MW96] for more kinds of analysis
involving the second approach.
We begin in Section 2 by introducing the necessary notation. Also discussed in this
section is the model of computation we adopt, its limitations, and the reasons we chose it.
One might expect that there are pairs of search algorithms A and B such that A performs better than B on average, even if B sometimes outperforms A. As an example, one
might expect that hill-climbing usually outperforms hill-descending if one's goal is to nd a
maximum of the cost function. One might also expect it would outperform a random search
in such a context.
One of the main results of this paper is that such expectations are incorrect. We prove
two NFL theorems in Section 3 that demonstrate this and more generally illuminate the
connection between algorithms and problems. Roughly speaking, we show that for both
static and time dependent optimization problems, the average performance of any pair of
algorithms across all possible problems is exactly identical. This means in particular that if
some algorithm a1's performance is superior to that of another algorithm a2 over some set of
optimization problems, then the reverse must be true over the set of all other optimization
problems. (The reader is urged to read this section carefully for a precise statement of these
theorems.) This is true even if one of the algorithms is random any algorithm a1 performs
worse than randomly just as readily (over the set of all optimization problems) as it performs
better than randomly. Possible objections to these results are also addressed in Sections 3.1
and 3.2.
In Section 4 we present a geometric interpretation of the NFL theorems. In particular,
2
we show that an algorithm's average performance is determined by how \aligned" it is with
the underlying probability distribution over optimization problems on which it is run. This
Section is critical for anyone wishing to understand how the NFL results are consistent with
the well-accepted fact that many search algorithms that do not take into account knowledge
concerning the cost function work quite well in practice
Section 5.1 demonstrates that the NFL theorems allow one to answer a number of what
would otherwise seem to be intractable questions. The implications of these answers for
measures of algorithm performance and of how best to compare optimization algorithms are
explored in Section 5.2.
In Section 6 we discuss some of the ways in which, despite the NFL theorems, algorithms can have a priori distinctions that hold even if nothing is specied concerning the
optimization problems. In particular, we show that there can be \head-to-head" minimax
distinctions between a pair of algorithms, it i.e., we show that considered one f at a time, a
pair of algorithms may be distinguishable, even if they are not when one looks over all f 's.
In Section 7 we present an introduction to the alternative approach to the formal analysis
of optimization in which problems are held xed and one looks at properties across the space
of algorithms. Since these results hold in general, they hold for any and all optimization
problems, and in this are independent of the what kinds of problems one is more or less likely
to encounter in the real world. In particular, these results state that one has no a priori
justication for using a search algorithm's behavior so far on a particular cost function
to predict its future behavior on that function. In fact when choosing between algorithms
based on their observed performance it does not suce to make an assumption about the cost
function some (currently poorly understood) assumptions are also being made about how
the algorithms in question are related to each other and to the cost function. In addition to
presenting results not found in MW96], this section serves as an introduction to perspective
adopted in MW96].
We conclude in Section 8 with a brief discussion, a summary of results, and a short list
of open problems.
We have conned as many of our proofs to appendices as possible to facilitate the ow
of the paper. A more detailed | and substantially longer | version of this paper, a version
that also analyzes some issues not addressed in this paper, can be found in WM95].
Finally, we cannot emphasize enough that no claims whatsoever are being made in
this paper concerning how well various search algorithms work in practice. The focus of
this paper is on what can be said a priori, without any assumptions and from mathematical
principles alone, concerning the utility of a search algorithm.
2 Preliminaries
We restrict attention to combinatorial optimization in which the search space, X , though
perhaps quite large, is nite. We further assume that the space of possible \cost" values, Y ,
is also nite. These restrictions are automatically met for optimization algorithms run on
digital computers. For example, typically Y is some 32 or 64 bit representation of the real
3
numbers in such a case.
The size of the spaces X and Y are indicated by jXj and jYj respectively. Optimization
problems f (sometimes called \cost functions" or \objective functions" or \energy functions") are represented as mappings f : X 7! Y . F = Y X is then the space of all possible
problems. F is of size jYjjXj | a very large but nite number. In addition to static f , we
shall also be interested in optimization problems that depend explicitly on time. The extra
notation needed for such time-dependent problems will be introduced as needed.
It is common in the optimization community to adopt an oracle-based view of computation. In this view, when assessing the performance of algorithms, results are stated in terms
of the number of function evaluations required to nd a certain solution. Unfortunately
though, many optimization algorithms are wasteful of function evaluations. In particular,
many algorithms do not remember where they have already searched and therefore often
revisit the same points. Although any algorithm that is wasteful in this fashion can be made
more ecient simply by remembering where it has been (c.f. tabu search Glo89, Glo90]),
many real-world algorithms elect not to employ this stratagem. Accordingly, from the point
of view of the oracle-based performance measures, there are \artefacts" distorting the apparent relationship between many such real-world algorithms.
This diculty is exacerbated by the fact that the amount of revisiting that occurs is
a complicated function of both the algorithm and the optimization problem, and therefore
cannot be simply \ltered out" of a mathematical analysis. Accordingly, we have elected to
circumvent the problem entirely by comparing algorithms based on the number of distinct
function evaluations they have performed. Note that this does not mean that we cannot
compare algorithms that are wasteful of evaluations | it simply means that we compare
algorithms by counting only their number of distinct calls to the oracle.
We call a time-ordered set of m distinct visited points a \sample" of size m. Samples are
denoted by dm f(dxm(1) dym (1)) (dxm(m) dym(m))g. The points in a sample are ordered
according to the time at which they were generated. Thus dxm (i) indicates the X value of
the ith successive element in a sample of size m and dym (i) is the associated cost or Y value.
dym fdym (1) dym (m)g will be used to indicate the ordered set of cost values. The space
of all samples of size m is Dm = (X Y )m (so dm 2 Dm ) and the set of all possible samples
of arbitrary size is D m 0Dm .
As an important clarication of this denition, consider a hill-descending algorithm.
This is the algorithm that examines a set of neighboring points in X and moves to the one
having the lowest cost. The process is then iterated from the newly chosen point. (Often,
implementations of hill-descending stop when they reach a local minimum, but they can
easily be extended to run longer by randomly jumping to a new unvisited point once the
neighborhood of a local minimum has been exhausted.) The point to note is that because
a sample contains all the previous points at which the oracles was consulted, it includes the
(X Y ) values of all the neighbors of the current point, and not only the lowest cost one that
the algorithm moves to. This must be taken into account when counting the value of m.
Optimization algorithms a are represented as mappings from previously visited sets of
points to a single new (i.e., previously unvisited) point in X . Formally, a : d 2 D 7!
fxjx 62 dX g. Given our decision to only measure distinct function evaluations even if an
4
algorithm revisits previously searched points, our denition of an algorithm includes all
common black-box optimization techniques like simulated annealing and evolutionary algorithms. (Techniques like branch and bound LW66] are not included since they rely explicitly
on the cost structure of partial solutions, and we are here interested primarily in black-box
algorithms.)
As dened above, a search algorithm is deterministic every sample maps to a unique new
point. Of course essentially all algorithms implemented on computers are deterministic1, and
in this our denition is not restrictive. Nonetheless, it is worth noting that all of our results
are extensible to non-deterministic algorithms, where the new point is chosen stochastically
from the set of unvisited points. (This point is returned to below.)
Under the oracle-based model of computation any measure of the performance of an
algorithm after m iterations is a function of the sample dym . Such performance measures
will be indicated by (dym ). As an example, if we are trying to nd a minimum of f , then
a reasonable measure of the performance of a might be the value of the lowest Y value in
dym : (dym ) = minifdym(i) : i = 1 : : : mg. Note that measures of performance based on factors
other than dym (e.g., wall clock time) are outside the scope of our results.
We shall cast all of our results in terms of probability theory. We do so for three reasons.
First, it allows simple generalization of our results to stochastic algorithms. Second, even
when the setting is deterministic, probability theory provides a simple consistent framework
in which to carry out proofs.
The third reason for using probability theory is perhaps the most interesting. A crucial
factor in the probabilistic framework is the distribution P (f ) = P (f (x1) f (xjXj)). This
distribution, dened over F , gives the probability that each f 2 F is the actual optimization
problem at hand. An approach based on this distribution has the immediate advantage that
often knowledge of a problem is statistical in nature and this information may be easily
encodable in P (f ). For example, Markov or Gibbs random eld descriptions KS80] of
families of optimization problems express P (f ) exactly.
However exploiting P (f ) also has advantages even when we are presented with a single
uniquely specied cost function. One such advantage is the fact that although it may be
fully specied, many aspects of the cost function are eectively unknown (e:g:, we certainly
do not know the extrema of the function.) It is in many ways most appropriate to have this
eective ignorance reected in the analysis as a probability distribution. More generally,
we usually act as though the cost function is partially unknown. For example, we might
use the same search algorithm for all cost functions in a class (e.g., all traveling salesman
problems having certain characteristics). In so doing, we are implicitly acknowledging that
we consider distinctions between the cost functions in that class to be irrelevant or at least
unexploitable. In this sense, even though we are presented with a single particular problem
from that class, we act as though we are presented with a probability distribution over cost
functions, a distribution that is non-zero only for members of that class of cost functions.
P (f ) is thus a prior specication of the class of the optimization problem at hand, with
dierent classes of problems corresponding to dierent choices of what algorithms we will
1
In particular, note that random number generators are deterministic given a seed.
5
use, and giving rise to dierent distributions P (f ).
Given our choice to use probability theory, the performance of an algorithm a iterated
m times on a cost function f is measured with P (dym jf m a). This is the conditional probability of obtaining a particular sample dm under the stated conditions. From P (dym jf m a)
performance measures (dym ) can be found easily.
In the next section we will analyze P (dym jf m a), and in particular how it can vary with
the algorithm a. Before proceeding with that analysis however, it is worth briey noting
that there are other formal approaches to the issues investigated in this paper. Perhaps the
most prominent of these is the eld of computational complexity. Unlike the approach taken
in this paper, computational complexity mostly ignores the statistical nature of search, and
concentrates instead on computational issues. Much (though by no means all) of computational complexity is concerned with physically unrealizable computational devices (Turing
machines) and the worst case amount of resources they require to nd optimal solutions. In
contrast, the analysis in this paper does not concern itself with the computational engine
used by the search algorithm, but rather concentrates exclusively on the underlying statistical nature of the search problem. In this the current probabilistic approach is complimentary
to computational complexity. Future work involves combining our analysis of the statistical
nature of search with practical concerns for computational resources.
3 The NFL theorems
In this section we analyze the connection between algorithms and cost functions. We have
dubbed the associated results \No Free Lunch" (NFL) theorems because they demonstrate
that if an algorithm performs well on a certain class of problems then it necessarily pays
for that with degraded performance on the set of all remaining problems. Additionally, the
name emphasizes the parallel with similar results in supervised learning Wol96a, Wol96b].
The precise question addressed in this section is: \How does the set of problems F1 F
for which algorithm a1 performs better than algorithm a2 compare to the set F2 F for
which the reverse is true?" To address this question we compare the sum over all f of
P (dym jf m a1) to the sum over all f of P (dym jf m a2). This comparison constitutes a major
result of this paper: P (dymjf m a) is independent of a when we average over all cost functions:
Theorem 1 For any pair of algorithms a1 and a2,
X
X
P (dym jf m a1) = P (dym jf m a2):
f
f
A proof of this result is found in Appendix A. An immediate corollary of this result is that for
any performance measure (dym ), the average over all f of P ((dym )jf m a) is independent
of a. The precise way that the sample is mapped to a performance measure is unimportant.
This theorem explicitly demonstrates that what an algorithm gains in performance on
one class of problems it necessarily pays for on the remaining problems that is the only way
that all algorithms can have the same f -averaged performance.
6
A result analogous to Theorem 1 holds for a class of time-dependent cost functions. The
time-dependent functions we consider begin with an initial cost function f1 that is present
at the sampling of the rst x value. Before the beginning of each subsequent iteration of
the optimization algorithm, the cost function is deformed to a new function, as specied
by a mapping T : F N ! F .2 We indicate this mapping with the notation Ti. So the
function present during the ith iteration is fi+1 = Ti(fi). Ti is assumed to be a (potentially
i-dependent) bijection between F and F . We impose bijectivity because if it did not hold,
the evolution of cost functions could narrow in on a region of f 's for which some algorithms
may perform better than others. This would constitute an a priori bias in favor of those
algorithms, a bias whose analysis we wish to defer to future work.
How best to assess the quality of an algorithm's performance on time-dependent cost
functions is not clear. Here we consider two schemes based on manipulations of the denition
of the sample. In scheme 1 the particular Y value in dym(j ) corresponding to a particular
x value dxm (j ) is given by the cost function that was present when dxm (j ) was sampled. In
contrast, for scheme 2 we imagine a sample Dmy given by the Y values from the present
cost function for each of the x values in dxm . Formally if dxm = fdxm (1) dxm(m)g, then
in scheme 1 we have dym = ff1(dxm(1)) Tm;1(fm;1)(dxm (m))g, and in scheme 2 we have
Dmy = ffm (dxm(1)) fm(dxm (m))g where fm = Tm;1(fm;1 ) is the nal cost function.
In some situations it may be that the members of the sample \live" for a long time, on
the time scale of the evolution of the cost function. In such situations it may be appropriate
to judge the quality of the search algorithm by Dmy all those previous elements of the sample
are still \alive" at time m, and therefore their current cost is of interest. On the other hand,
if members of the sample live for only a short time on the time scale of evolution of the cost
function, one may instead be concerned with things like how well the \living" member(s) of
the sample track the changing cost function. In such situations, it may make more sense to
judge the quality of the algorithm with the dym sample.
Results similar to Theorem 1 can be derived for both schemes. By analogy with that
theorem, we average over all possible ways a cost function may
be time-dependent, i.e., we
P
average over all T (rather than over all f ). Thus we consider T P (dym jf1 T m a) where f1
is the initial cost function. Since T only takes eect for m > 1, and since f1 is xed, there
are a priori distinctions between algorithms as far as the rst member of the population is
concerned. However after redening samples to only contain those elements added after the
rst iteration of the algorithm, we arrive at the following result, proven in Appendix B:
Theorem 2 For all dym, Dmy , m > 1, algorithms a1 and a2, and initial cost functions f1,
X
X
P (dym jf1 T m a1) = P (dym jf1 T m a2):
T
and
X
T
T
P (Dmy jf1 T m a1) =
X
T
P (Dmy jf1 T m a2):
An obvious restriction would be to require that doesn't vary with time, so that it is a mapping simply
from F to F . An analysis for 's limited this way is beyond the scope of this paper.
2
T
T
7
So in particular, if one algorithm outperforms another for certain kinds of evolution operators,
then the reverse must be true on the set of all other evolution operators.
Although this particular result is similar to the NFL result for the static case, in general
the time-dependent situation is more subtle. In particular, with time-dependence there are
situations in which there can be a priori distinctions between algorithms even for those
members of the population arising after the rst. For example,
in general there will be
distinctions between algorithms when considering the quantity Pf P (dym jf T m a). To see
this, consider the case where X is a set of contiguous integers and for all iterations T is a
shift operator, replacing f (x) by f (x ; 1) for all x (with min(x) ; 1 max(x)). For such a
case we can construct algorithms which behave dierently a priori. For example, take a to
be the algorithm that rst samples f at x1, next at x1 + 1 and so on, regardless of the values
in
the population. Then for any f , dym is always made up of identical Y values. Accordingly,
P
y
y
y
f P (dm jf T m a) is non-zero only for dm for which all values dm (i) are identical. Other
search algorithms, even for the same shift T , do not have this restriction on Y values. This
constitutes an a priori distinction between algorithms.
3.1 Implications of the NFL theorems
As emphasized above, the NFL theorems mean that if an algorithm does particularly well on
one class of problems then it most do more poorly over the remaining problems. In particular,
if an algorithm performs better than random search on some class of problems then in must
perform worse than random search on the remaining problems. Thus comparisons reporting
the performance of a particular algorithm with particular parameter setting on a few sample
problems are of limited utility. While sicj results do indicate behavior on the narrow range
of problems considered, one should be very wary of trying to generalize those results to other
problems.
Note though that the NFL theorem need not be viewed this way, as a way of comparing
function classes F1 and F2 (or classes of evolution operators T1 and T2, as the case might
be). It can be viewed instead as a statement concerning any algorithm's performance when
f is not xed, under the uniform prior over cost functions, P (f ) = 1=jFj. If we wish instead
to analyze performance where f is not xed, as in this alternative interprations of the NFL
theorem, but in contrast with the NFL case f is now chosen from a non-uniform prior, then
we must analyze explicitly the sum
P (dym jm a) =
X
f
P (dym jf m a)P (f )
(1)
Since it is certainly true that any class of problems faced by a practitioner will not have a at
prior, what are the practical implications of the NFL theorems when viewed as a statement
concerning an algorithm's performance for non-xed f ? This question is taken up in greater
detail in Section 4 but we make a few comments here.
First, if the practitioner has knowledge of problem characteristics but does not incorporate them into the optimization algorithm, then P (f ) is eectively uniform. (Recall that
8
P (f ) can be viewed as a statement concerning the practitioner's choice of optimization algorithms.) In such a case, the NFL theorems establish that there are no formal assurances
that the algorithm chosen will be at all eective.
Secondly, while most classes of problems will certainly have some structure which, if
known, might be exploitable, the simple existence of that structure does not justify choice of
a particular algorithm that structure must be known and reected directly in the choice of
algorithm to serve as such a justication. In other words, the simple existence of structure
per se, absent a specication of that structure, cannot provide a basis for preferring one algorithm over another. Formally, this is established by the existence of NFL-type theorems in
which rather than average over specic cost functions f , one averages over specic \kinds of
structure", i.e., theorems in which one averages P (dym j m a) over distributions P (f ). That
such theorems hold when one averages over all P (f ) means that the indistinguishability of
algorithms associated with uniform P (f ) is not some pathological, outlier case. Rather uniform P (f ) is a \typical" distribution as far as indistinguishability of algorithms is concerned.
The simple fact that the P (f ) at hand is non-uniform cannot serve to determine one's choice
of optimization algorithm.
Finally, it is important to emphasize that even if one is considering the case where f is
not xed, performing the associated average according to a uniform P (f ) is not essential for
NFL to hold. NFL can Qalso be demonstrated for a range of non-uniform priors. For example,
any prior of the form x2X P 0(f (x)) (where P 0(y = f (x)) is the distribution of Y values)
will also give NFL. The f -average can also enforce correlations between costs at dierent
X values and NFL still obtain. For example if costs are rank ordered (with ties broken in
some arbitrary way) and we sum only over all cost functions given by permutations of those
orders, then NFL still holds.
The choice of uniform P (f ) was motivated more from theoretical rather pragramattic
concerns, as a way of analyzing the theoretical structure of optimization. Nevertheless, the
cautionary observations presented above make clear that an analysis of the uniform P (f )
case has a number of ramications for practitioners.
3.2 Stochastic optimization algorithms
Thus far we have considered the case in which algorithms are deterministic. What is the situation for stochastic algorithms? As it turns out, NFL results hold even for such algorithms.
The proof of this is straightforward. Let be a stochastic \non-potentially revisiting"
algorithm. Formally, this means that is a mapping taking any d to a d-dependent distribution over X that equals zero for all x 2 dx. (In this sense is what in statistics community is
known as a \hyper-parameter", specifying the function P (dxm+1 (m + 1) j dm ) for all m and
d.) One can now reproduce the derivation of the NFL result for deterministic algorithms,
only with a replaced by throughout. In so doing all steps in the proof remain valid. This
establishes that NFL results apply to stochastic algorithms as well as deterministic ones.
9
4 A geometric perspective on the NFL theorems
Intuitively, the NFL theorem illustrates that even if knowledge of f (perhaps specied
through P (f )) is not incorporated into a, then there are no formal assurances that a will
be eective. Rather, eective optimization relies on a fortuitous matching between f and a.
This point is formally established by viewing the NFL theorem from a geometric perspective.
Consider the space F of all possible cost functions. As previously discussed in regard to
Equation (1) the probability of obtaining some dym is
P (dym jm a) =
X
f
P (dym jm a f )P (f ):
where P (f ) is the prior probability that the optimization problem at hand has cost function
f . This sum over functions can be viewed as an inner product in F . More precisely, dening
the F -space vectors ~vdym a m and ~p by their f components ~vdym a m(f ) P (dym jm a f ) and
~p(f ) P (f ) respectively,
P (dym jm a) = ~vdym a m p~:
(2)
This equation provides a geometric interpretation of the optimization process. dym can
be viewed as xed to the sample that is desired, usually one with a low cost value, and m
is a measure of the computational resources that can be aorded. Any knowledge of the
properties of the cost function goes into the prior over cost functions, p~. Then Equation
(2) says the performance of an algorithm is determined by the magnitude of its projection
onto ~p, i.e. by how aligned ~vdym a m is with the problems ~p. Alternatively, by averaging over
dym , it is easy to see that E (dym jm a) is an inner product between ~p and E (dym jm a f ). The
expectation of any performance measure (dym) can be written similarly.
In any of these cases, P (f ) or ~p must \match" or be aligned with a to get desired
behavior. This need for matching provides a new perspective on how certain algorithms can
perform well in practice on specic kinds of problems. For example, it means that the years
of research into the traveling salesman problem (TSP) have resulted in algorithms aligned
with the (implicit) ~p describing traveling salesman P
problems of interest to TSP researchers.
Taking the geometric view, the NFL result that f P (dym jf m a) is independent of a has
the interpretation that for any particular dym and m, all algorithms a have the same projection
onto the the uniform P (f ), represented by the diagonal vector ~1. Formally, vdym a m ~1 =
cst(dym m). For deterministic algorithms the components of vdym a m (i.e. the probabilities
that algorithm a gives sample dym on cost function
f after m distinct cost evaluations) are
P 2 y
all either 0 or 1, so NFL also implies that f P (dm j m a f ) = cst(dym m). Geometrically,
this indicates that the length of ~vdym a m is independent of a. Dierent algorithms thus
generate dierent vectors ~vdym a m all having the same length and lying on a cone with constant
projection onto ~1. (A schematic of this situation is shown in Figure 1 for the case where
F is 3-dimensional.) Because the components of ~vc a m are binary we might equivalently
view ~vdym a m as lying on the subset the vertices of the Boolean hypercube having the same
hamming distance from ~1.
10
1
p
Figure 1: Schematic view of the situation in which function space F is 3-dimensional. The
uniform prior over this space, ~1 lies along the diagonal. Dierent algorithms a give dierent
vectors v lying in the cone surrounding the diagonal. A particular problem is represented by
its prior ~p lying on the simplex. The algorithm that will perform best will be the algorithm
in the cone having the largest inner product with p~.
Now restrict attention to algorithms having the same probability of some particular dym.
The algorithms in this set lie in the intersection of 2 cones|one about the diagonal, set by
the NFL theorem, and one set by having the same probability for dym. This is in general an
jFj; 2 dimensional manifold. Continuing, as we impose yet more dym -based restrictions on a
set of algorithms, we will continue to reduce the dimensionality of the manifold by focusing
on intersections of more and more cones.
The geometric view of optimization also suggests alternative measures for determining
how \similar" two optimization algorithms are. Consider again Equation (2). In that the
algorithm directly only gives ~vdym a m, perhaps the most straight-forward way to compare two
algorithms a1 and a2 would be by measuring how similar the vectors ~vdym a1 m and ~vdym a2 m are.
(E.g., by evaluating the dot product of those vectors.) However those vectors occur on the
right-hand side of Equation (2), whereas the performance of the algorithms | which is after
all our ultimate concern | instead occur on the left-hand side. This suggests measuring
the similarity of two algorithms not directly in terms of their vectors ~vdym a m, but rather in
terms of the dot products of those vectors with ~p. For example, it may be the case that
algorithms behave very similarly for certain P (f ) but are quite dierent for other P (f ). In
many respects, knowing this about two algorithms is of more interest than knowing how
their vectors ~vdym a m compare.
As another example of a similarity measure suggested by the geometric perspective,
we could measure similarity between algorithms based on similarities between P (f )'s. For
example, for two dierent algorithms, one can imagine solving for the P (f ) that optimizes
11
P (dym j m a) for those algorithms, in some non-trivial sense.3 We could then use some
measure of distance between those two P (f ) distributions as a gauge of how similar the
associated algorithms are.
Unfortunately, exploiting the inner product formula in practice, by going from a P (f )
to an algorithm optimal for that P (f ), appears to often be quite dicult. Indeed, even
determining a plausible P (f ) for the situation at hand is often dicult. Consider, for
example, TSP problems with N cities. To the degree that any practitioner attacks all
N -city TSP cost functions with the same algorithm, that practitioner implicitly ignores
distinctions between such cost functions. In this, that practitioner has implicitly agreed
that the problem is one of how their xed algorithm does across the set of all N -city TSP
cost functions. However the detailed nature of the P (f ) that is uniform over this class of
problems appears to be dicult to elucidate.
On the other hand, there is a growing body of work that does rely explicitly on enumeration of P (f ). For example, applications of Markov random elds Gri76, KS80] to cost
landscapes yield P (f ) directly as a Gibbs distribution.
5 Calculational applications of the NFL theorems
In this section we explore some of the applications of the NFL theorems for performing
calculations concerning optimization. We will consider both calculations of practical and
theoretical interest, and begin with calculations of theoretical interest, in which informationtheoretic quantities arise naturally.
5.1 Information-theoretic aspects of optimization
For expository purposes, we simplify the discussion slightly by considering only the histogram
of number of instances of each possible cost value produced by a run of an algorithm, and
not the temporal order in which those cost values were generated. (Essentially all realworld performance measures are independent of such temporal information.) We indicate
that histogram with the symbol ~c ~c has Y components (cY1 cY2 cY ), where ci is the
number of times cost value Yi occurs in the sample dym .
Now consider any question like the following: \What fraction of cost functions give a
particular histogram ~c of cost values after m distinct cost evaluations produced by using a
particular instantiation of an evolutionary algorithm FOW66, Hol93]?"
At rst glance this seems to be an intractable question. However it turn out that the
NFL theorem provides a way to answer it. This is because | according to the NFL theorem
| the answer must be independent of the algorithm used to generate ~c. Consequently we
can chose an algorithm for which the calculation is tractable.
jY j
In particular, one may want to impose restrictions on ( ). For instance, one may wish to only consider
( ) that are invariant under at least partial relabelling of the elements in X , to preclude there being an
algorithm that will assuredly \luck out" and land on minx2X ( ) on its very rst query.
3
P f
P f
f x
12
Theorem 3 For any algorithm, the fraction of cost functions that result in a particular
histogram ~c = m~ is
f (~ ) =
m
c1 c2 c
jY j
jYjjXj;m
jYjjXj
=
m
c1 c2 c
jYjm
jY j
:
For large enough m this can be approximated as
exp m S (~ )]
f (~ ) = C (m jYj) QjYj 1=2
i=1 i
where S (~ ) is the entropy of the distribution ~ , and C (m jYj) is a constant that does not
depend on ~ .
This theorem is derived in Appendix C. If some of the ~ i are 0, the approximation still holds,
only with Y redened to exclude the y's corresponding to the zero-valued ~ i . However Y
is dened, the normalization constant of Equation (3) can be found by summing over all ~
lying on the unit simplex ?].
A question related to one addressed in this theorem is the following: \For a given cost
function, what is the fraction alg of all algorithms that give rise to a particular ~c?" It turns
out that the only feature of f relevant for this question is the histogram of its cost values
formed by looking across all X . Specify the fractional form of this histogram by ~ : there
are Ni = i jXj points in X for which f (x) has the i'th Y value.
In Appendix D it is shown that to leading order, alg(~ ~ ) depends on yet another
information theoretic quantity, the Kullback-Liebler distance CT91] between ~ and ~ :
Theorem 4 For a given f with histogram N~ = jXj~ , the fraction of algorithms that give
rise to a histogram ~c = m~ is given by
alg(~ ~ ) =
QjYj Ni i=1 ci
jXj :
m
For large enough m this can be written as
;m D (~ ~)
e KL
alg(~ ~ ) = C (m jXj jYj) QjYj 1=2
i=1 i
where DKL (~ ~ ) is the Kullback-Liebler distnace between the distributions and .
As before, C can be calculated by summing ~ over the unit simplex.
13
(3)
5.2 Measures of performance
We now show how to apply the NFL framework to calculate certain benchmark performance
measures. These allow both the programmatic (rather than ad hoc) assessment of the ecacy
of any individual optimization algorithm and principled comparisons between algorithms.
Without loss of generality, assume that the goal of the search process is nding a minimum. So we are interested in the -dependence of P (min(~c) > j f m a), by which we mean
the probability that the minimum cost an algorithm a nds on problem f in m distinct
evaluations is larger than . At least three quantities related to this conditional probability
can be used to gauge an algorithm's performance in a particular optimization run:
i) The uniform average of P (min(~c) > j f m a) over all cost functions.
ii) The form P (min(~c) > j f m a) takes for the random algorithm, which uses no information from the sample dm .
iii) The fraction of algorithms which, for a particular f and m, result in a ~c whose minimum
exceeds .
These measures give benchmarks which any algorithm run on a particular cost function
should surpass if that algorithm is to be considered as having worked well for that cost
function.
Without loss of generality assume the that i'th cost value (i.e., Yi equals i. So cost values
run from a minimum of 1 to a maximum of jYj, in integer increments. The following results
are derived in Appendix E.
Theorem 5
X
f
P (min(~c) > j f m) = !m()
where ! () 1 ; =jYj is the fraction of cost lying above . In the limit of jYj ! 1, this
distribution obeys the following relationship
P
f E (min(~c) j f
jYj
m) = 1 :
m+1
Unless one's algorithm has its best-cost-so-far drop faster than the drop associated with
these results, one would be hard-pressed indeed to claim that the algorithm is well-suited to
the cost function at hand. After all, for such performance the algorithm is doing no better
than one would expect it to for a randomly chosen cost function.
Unlike the preceding measure, the measures analyzed below take into account the actual
cost function at hand. This is manifested in the dependance of the values of those measures
on the vector N~ given by the cost function's histogram (N~ = jXj~ ):
14
Theorem 6 For the random algorithm ~a,
P (min(~c) j f m ~a) =
mY
;1
i=0
() ; i=jXj :
1 ; i=jXj
(4)
P
where () jYj
i= Ni =jXj is the fraction of points in X for which f (x) . To rst order
in 1=jXj
1
m
(
m
;
1)(1
;
(
))
(5)
P (min(~c) > j f m ~a) =
1;
2 ()
jXj + :
This result allows the calculation of other quantities of interest for measuring performance,
for example the quantity
m ()
E (min(~c)jf m ~a) =
jYj
X
=1
P (min(~c) j f m ~a) ; P (min(~c) + 1 j f m ~a)]:
Note that for many cost functions of both practical and theoretical interest, cost values are
distributed Gaussianly. For such cases, we can use that Gaussian nature of the distribution
to facilitate our calculations. In particular, if the mean
p and variance of the Gaussian are and respectively, then we have () = erfc((;)= 2)=2, where erfc is the complimentary
error function.
To calculate the third performance measure, note that for xed f and m, for any (deterministic) algorithm a, P (~c > j f m a) is either 1 or 0. Therefore the fraction of algorithms
which result in a ~c whose minimum exceeds is given by
P
c) > j f m a) :
a P (min(~P
a1
P
Expanding
in
terms
of
~
c
,
we
can
rewrite
the
numerator
of
this
ratio
as
~c P (min(~c) >
P
P
j~c) a P (~c j f m a). However the ratio of this quantity to a 1 is exactly what was calculated when we evaluated measure (ii) (see the beginning of the argument deriving Equation
(4)). This establishes the following:
Theorem 7 For xed f and m, the fraction of algorithms which result in a ~c whose minimum
exceeds is given by the quantity on the right-hand sides of Equations (4) and (5).
As a particular example of applying this result, consider measuring the value of min(~c)
produced in a particular run of your algorithm. Then imagine that when it is evaluated for
equal to this value, the quantity given in Equation (5) is less than 1=2. In such a situation
the algorithm in question has performaed worse than over half of all search algorithms, for
the f and m at hand hardly a stirring endorsement.
None of the discussion above explicitly concerns the dynamics of an algorithm's performance as m increases. Many aspects of such dynamics may be of interest. As an example, let
15
us consider whether, as m grows, there is any change in how well the algorithm's performance
compares to that of the random algorithm.
To this end, let the sample generated by the algorithm a after m steps be dm , and dene
0
y min(dym ). Let k be the number of additional steps it takes the algorithm to nd an
x such that f (x) < y0. Now we can estimate the number of steps it would have taken the
random search algorithm to search X ; dX and nd a point whose y was less than y0. The
expected value of this number of steps is 1=z(d) ; 1, where z(d) is the fraction of X ; dxm
for which f (x) < y0. Therefore k + 1 ; 1=z(d) is how much worse a did than would have the
random algorithm, on average.
Next imagine letting a run for many steps over some tness function f and plotting how
well a did in comparison to the random algorithm on that run, as m increased. Consider
the step where a nds its n'th new value of min(~c). For that step, there is an associated k
(the number of steps until the next min(dym )) and z(d). Accordingly, indicate that step on
our plot as the point (n k + 1 ; 1=z(d)). Put down as many points on our plot as there are
successive values of min(~c(d)) in the run of a over f .
If throughout the run a is always a better match to f than is the random search algorithm,
then all the points in the plot will have their ordinate values lie below 0. If the random
algorithm won for any of the comparisons though, that would mean a point lying above 0.
In general, even if the points all lie to one side of 0, one would expect that as the search
progresses there is corresponding (perhaps systematic) variation in how far away from 0 the
points lie. That variation tells one when the algorithm is entering harder or easier parts of
the search.
Note that even for a xed f , by using dierent starting points for the algorithm one
could generate many of these plots and then superimpose them. This allows a plot of
the mean value of k + 1 ; 1=z(d) as a function of n along with an associated error bar.
Similarly, one could replace the single number z(d) characterizing the random algorithm
with a full distribution over the number of required steps to nd a new minimum. In these
and similar ways, one can generate a more nuanced picture of an algorithm's performance
than is provided by any of the single numbers given by the performance measure discussed
above.
6 Minimax distinctions between algorithms
The NFL theorems do not direclty address minimax properties of search. For example, say
we're considering two deterministic algorithms, a1 and a2. It may very well be that there
exist cost functions f such that a1's histogram is much better (according to some appropriate
performance measure) than a2's, but no cost functions for which the reverse is true. For the
NFL theorem to be obeyed in such a scenario, it would have to be true that there are many
more f for which a2's histogram is better than a1's than vice-versa, but it is only slightly
better for all those f . For such a scenario, in a certain sense a1 has better \head-to-head"
minimax behavior than a2 there are f for which a1 beats a2 badly, but none for which a1
does substantially worse than a2.
16
Formally, we say that there exists head-to-head minimax distinctions between two algorithms a1 and a2 i there exists a k such that for at least one cost function f , the dierence
E (~c j f m a1) ; E (~c j f m a2) = k, but there is no other f for which E (~c j f m a2) ; E (~c j
f m a1) = k. (A similar denition can be used if one is instead interested in (~c) or dym
rather than ~c.)
It appears that analyzing head-to-head minimax properties of algorithms is substantially
more dicult than analyzing average behavior (like in the NFL theorem). Presently, very
little is known about minimax behavior involving stochastic algorithms. In particular, it is
not known if there are any senses in which a stochastic version of a deterministic algorithm
has better/worse minimax behavior than that deterministic algorithm. In fact, even if we
stick completely to deterministic algorithms, only an extremely preliminary understanding
of minimax issues has been reached.
What we do know is the following. Consider the quantity
X
f
Pdym 1 dym 2 (z z0 j f m a1 a2)
for deterministic algorithms a1 and a2. (By PA (a) is meant the distribution of a random
variable A evaluated at A = a.) For deterministic algorithms, this quantity is just the
number of f such that it is both true that a1 produces a population with Y components z
and that a2 produces a population with Y components z0.
In Appendix F, it is proven by example that this quantity need not be symmetric under
interchange of z and z0:
Theorem 8 In general,
X
f
Pdym 1 dym 2 (z z0 j f m a1 a2) 6=
X
f
Pdym 1 dym 2 (z0 z j f m a1 a2):
(6)
This means that under certain circumstances, even knowing only the Y components of the
populations produced by two algorithms run on the same (unknown) f , we can infer something concerning what algorithm produced each population.
Now consider the quantity
X
f
PC1 C2 (z z0 j f m a1 a2)
again for deterministic algorithms a1 and a2. This quantity is just the number of f such that
it is both true that a1 produces a histogram z and that a2 produces a histogram z0. It too
need not be symmetric under interchange of z and z0 (see Appendix F). This is a stronger
statement then the asymmetry of dy 's statement, since any particular histogram corresponds
to multiple populations.
It would seem that neither of these two results directly implies that there are algorithms
a1 and a2 such that for some f a1's histogram is much better than a2's, but for no f 's is the
reverse is true. To investigate this problem involves looking over all pairs of histograms (one
17
pair for each f ) such that there is the same relationship between (the performances of the
algorithms, as reected in) the histograms. Simply having an inequality between the sums
presented above does not seem to directly imply that the relative performances between
the associated pair of histograms is asymmetric. (To formally establish this would involve
creating scenarios in which there is an inequality between the sums, but no head-to-head
minimax distinctions. Such an analysis is beyond the scope of this paper.)
On the other hand, having the sums equal does carry obvious implications for whether
there are head-to-head minimax distinctions. For example, if both algorithms are deterministic, then for any particular f PPdym 1 dym 2 (z1 z2 j f m a1 a2) equals 1 for one (z1 z2) pair, and 0
for all others. In such a case,P f Pdym 1 dym 2 (z1 z2 j f m a1 a2) P
is just the number of f that re0
y
y
sult in the pair (z1 z2). So f Pdm 1 dm 2 (z z j f m a1 a2) = f Pdym 1 dym 2 (z0 z j f m a1 a2)
implies that there are no head-to-head minimax distinctions between a1 and a2. The converse
does not appear to hold however.4
As a preliminary analysis of whether there can be head-to-head minimax distinctions, we
can exploit the result in Appendix F, which concerns the case where jXj = jYj = 3. First,
dene the following performance measures of two-element populations, Q(dy2 ):
i) Q(y2 y3) = Q(y3 y2) = 2.
ii) Q(y1 y2) = Q(y2 y1) = 0.
iii) Q of any other argument = 1.
In Appendix F we show that for this scenario there exist pairs of algorithms a1 and a2 such
that for one f a1 generates the histogram fy1 y2g and a2 generates the histogram fy2 y3g,
but there is no f for which the reverse occurs (i.e., there is no f such that a1 generates the
histogram fy2 y3g and a2 generates fy1 y2g).
So in this scenario, with our dened performance measure, there are minimax distinctions between a1 and a2. For one f the performance measures of algorithms a1 and a2 are
respectively 0 and 2. The dierence in the Q values for the two algorithms is 2 for that f .
However there are no other f for which the dierence is -2. For this Q then, algorithm a2 is
minimax superior to algorithm a1.
It is not currently known what restrictions on Q(dym) are needed for there to be minimax
distinctions between the algorithms. As an example, it may well be that for Q(dym ) =
minifdym (i)g there are no minimax distinctions between algorithms.
More generally, at present nothing is known about \how big a problem" these kinds of
asymmetries are. All of the examples of asymmetry considered here arise when the set of
Consider the grid of all ( 0) pairs. Assign to each grid point the number of that result in that grid
point's ( 0) pair. Then our constraints are i) by the hypothesis that there are no head-to-head minimax
distinctions, if grid point ( 1 2) is assigned a non-zero number, then so is ( 2 1 ) and ii) by the no-freelunch theorem, the sum of all numbers in row equals the sum of all numbers in column . These two
constraints do not appear to imply that the distribution of numbers is symmetric under interchange of rows
and columns. Although again, like before, to formally establish this point would involve explicitly creating
search scenarios in which it holds.
4
z z
f
z z
z
z
z
z
z
z
18
X values a1 has visited overlaps with those that a2 has visited. Given such overlap, and
certain properties of how the algorithms generated the overlap, asymmetry arises. A precise
specication of those \certain properties" is not yet in hand. Nor is it known how generic
they are, i.e., for what percentage of pairs of algorithms they arise. Although such issues are
easy to state (see Appendix F), it is not at all clear how best to answer them.
However consider the case where we are assured that in m steps the populations of two
particular algorithms have not overlapped. Such assurances hold, for example, if we are
comparing two hill-climbing algorithms that start far apart (on the scale of m) in X . It
turns out that given such assurances, there are no asymmetries between the two algorithms
for m-element populations. To see this formally, go throughP the argument used to prove
the NFL theorem, but apply that argument to the quantity f Pdym 1 dym 2 (z z0 j f m a1 a2)
rather than P (~c j f m a). Doing this establishes the following:
Theorem: If there is no overlap between dxm 1 and dxm 2, then
X
f
Pdym 1 dym 2 (z z0 j f m a1 a2) =
X
f
Pdym 1 dym 2 (z0 z j f m a1 a2):
(7)
An immediate
consequence of this theorem is that under the no-overlap conditions, the
P
quantity f PC1 C2 (z z0 j f m a1 a2) is symmetric under interchange of z and z0, as are
all distributions determined from this one over C1 and C2 (e.g., the distribution over the
dierence between those C 's extrema).
Note that with stochastic algorithms, if they give non-zero probability to all dxm, there
is always overlap to consider. So there is always the possibility of asymmetry between
algorithms if one of them is stochastic.
7
P (f )-independent
results
All work to this point has largely considered the behavior of various algorithms across a wide
range of problems. In this section we introduce the kinds of results that can be obtained
when we reverse roles and consider the properties of many algorithms on a single problem.
More results of this type are found in MW96]. The results of this section, although less
sweeping than the NFL results, hold no matter what the real world's distribution over cost
functions is.
Let a and a0 be two search algorithms. Dene a \choosing procedure" as a rule that
examines the samples dm and d0m , produced by a and a0 respectively, and based on those
populations, decides to use either a or a0 for the subsequent part of the search. As an
example, one \rational" choosing procedure is to use a for the subsequent part of the search
if and only it has generated a lower cost value in its sample than has a0. Conversely we
can consider a \irrational" choosing procedure that went with the algorithm that had not
generated the sample with the lowest cost solution.
At the point that a choosing procedure takes eect the cost function will have been
sampled at d dm d0m . Accordingly, if d>m refers to the samples of the cost function that
19
come after using the choosing algorithm, then the user is interested in the remaining sample
d>m . As always, without loss of generality it is assumed that the search algorithm chosen
by the choosing procedure does not return to any points in d.5
The following theorem, proven in Appendix G, establishes that there is no a priori
justication for using any particular choosing procedure. Loosely speaking, no matter what
the cost function, without special consideration of the algorithm at hand, simply observing
how well that algorithm has done so far tells us nothing a priori about how well it would do
if we continue to use it on the same cost function. For simplicity, in stating the result we
only consider deterministic algorithms.
Theorem 9 Let dm and d0m be two xed samples of size m, that are generated when the
algorithms a and a0 respectively are run on the (arbitrary) cost function at hand. Let A and
B be two dierent choosing procedures. Let k be the number of elements in c>m. Then
X
aa
0
X
P (c>m j f d d0 k a a0 A) =
aa
0
P (c>m j f d d0 k a a0 B ):
Implicit in this result is the assumption that the sum excludes those algorithms a and a0
that do not result in d and d0 respectively when run on f .
In the precise form it is presented above, the result may appear misleading, since it
treats all populations equally, when for any given f some populations will be more likely
than others. However even if one weights populations according to their probability of
occurrence, it is still true that, on average, the choosing procedure one uses has no eect on
likely c>m. This is established by the following result, proven in Appendix H:
Theorem 10 Under the conditions given in the preceding theorem,
X
X
P (c>m j f m k a a0 A) = P (c>m j f m k a a0 B ):
aa
aa
0
0
These results show that no assumption for P (f ) alone justies using some choosing
procedure as far as subsequent search is concerned. To have an intelligent choosing procedure,
one must take into account not only P (f ) but also the search algorithms one is choosing
among. This conclusion may be surprising. In particular, note that it means that there is no
intrinsic advantage to using a rational choosing procedure, which continues with the better
of a and a0, rather than using a irrational choosing procedure which does the opposite.
These results also have interesting implications for degenerate choosing procedures A falways use algorithm ag, and B falways use algorithm a0g. As applied to this case, they
can know to avoid the elements it has seen before. However a priori, has no way to avoid the elements
it hasn't seen yet but that 0 has (and vice-versa). Rather than have the denition of somehow depend
on the elements in 0 ; (and similarly for 0 ), we deal with this problem by dening >m to be set only by
those elements in >m that lie outside of . (This is similar to the convention we exploited above to deal
with potentially retracing algorithms.) Formally, this means that the random variable >m is a function of
as well as of >m . It also means there may be fewer elements in the histogram >m than there are in the
population >m .
5a
a
a
d
d
a
d
a
c
d
c
d
d
c
d
20
mean that for xed f1 and f2, if f1 does better (on average) with the algorithms in some set
A, then f2 does better (on average) with the algorithms in the set of all other algorithms.
In particular, if for some favorite algorithms a certain \well-behaved" f results in better
performance than does the random f , then that well-behaved f gives worse than random
behavior on the set all remaining algorithms. In this sense, just as there are no universally
ecacious search algorithms, there are no universally benign f which can be assured of
resulting in better than random performance regardless of one's algorithm.
In fact, things may very well be worse than this. In supervised learning, there is a
related result Wol96a]. Translated into the current context that result suggests that if one
restricts our sums to only be over those algorithms that are a good match to P (f ), then it is
often the case that\stupid" choosing procedures | like the irrational procedure of choosing
the algorithm with the less desirable ~c | outperform \intelligent" ones. What the set of
algorithms summed over must be for a rational choosing procedure to be superior to an
irrational is not currently known.
8 Conclusions
A framework has been presented in which to compare general-purpose optimization algorithms. A number of NFL theorems were derived that demonstrate the danger of comparing
algorithms by their performance on a small sample of problems. These same results also indicate the importance of incorporating problem-specic knowledge into the behavior of the
algorithm. A geometric interpretation was given showing what it means for an algorithm to
be well-suited to solving a certain class of problems. The geometric perspective also suggests
a number of measures to compare the similarity of various optimization algorithms.
More direct calculational applications of the NFL theorem were demonstrated by investigating certain information theoretic aspects of search, as well as by developing a number
of benchmark measures of algorithm performance. These benchmark measures should prove
useful in practice.
We provided an analysis of the ways that algorithms can dier a priori despite the
NFL theorems. We have also provided an introduction to a variant of the framework that
focuses on the behavior of a range of algorithms on specic problems (rather than specic
algorithms over a range of problems). This variant leads directly to reconsideration of many
issues addressed by computational complexity, as detailed in MW96].
Much future work clearly remains | the reader is directed to WM95] for a list of some
of it. Most important is the development of practical applications of these ideas. Can the geometric viewpoint be used to construct new optimization techniques in practice? We believe
the answer to be yes. At a minimum, as Markov random eld models of landscapes become
more wide-spread, the approach embodied in this paper should nd wider applicability.
Acknowledgments
We would like to thank Raja Das, David Fogel, Tal Grossman, Paul Helman, Bennett Levitan, Una-May O'Rielly and the reviewers for helpful comments and suggestions. WGM
thanks the Santa Fe Institute for funding and DHW thanks the Santa Fe Institute and TXN
21
Inc. for support.
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A NFL proof for static cost functions
We show that Pf P (~c j f m a) has no dependence on a. Conceptually, the proof is quite
simple but necessary book-keeping complicates things, lengthening the proof considerably.
The intuition behind the proof is quite simple though: by summing over all f we ensure that
22
the past performance of an algorithm has no bearing on its future performance. Accordingly,
under such a sum, all algorithms perform equally.
The proof is by induction. The induction is based on m = 1 and the inductive step is
based on breaking f into two independent parts, one for x 2 dxm and one for x 62 dxm . These
are evaluated separately, giving the desired result.
For m = 1 we write the sample as d1 = fdx1 f (dx1 )g where dx1 is set by a. The only
possible value for dy1 is f (dx1 ), so we have :
X
X
P (dy1 j f m = 1 a) = (dy1 f (dx1 ))
f
f
where is the Kronecker delta function.
Summing over all possible cost functions, (dy1 f (dx1 )) is 1 only for those functions which
have cost dy1 at point dx1 . Therefore that sum equals jYjjXj;1, independent of dx1 :
X
P (dy1 j f m = 1 a) = jYjjXj;1
f
which is independent of a. This bases the
induction.
P
y
y
The inductive
step
requires
that
if
P
f (dm jf m a) is independent of a for all dm , then
P
so also is f P (dym+1jf m + 1 a). Establishing this step completes the proof.
We begin by writing
P (dym+1 jf m + 1 a) = P (fdym+1 (1) : : : dym+1 (m)g dym+1(m + 1)jf m + 1 a)
= P (dym dym+1 (m + 1)jf m + 1 a)
= P (dym+1 (m + 1)jdm f m + 1 a) P (dym jf m + 1 a)
and thus
X
X
P (dym+1 jf m + 1 a) = P (dym+1(m + 1)jdym f m + 1 a) P (dym jf m + 1 a):
f
f
The new y value, dym+1(m + 1), will depend on the new x value, f and nothing else. So we
expand over these possible x values, obtaining
X
X
P (dym+1 jf m +1 a) = P (dym+1 (m + 1)jf x) P (xjdym f m +1 a)P (dymjf m + 1 a)
f
=
fx
X
fx
(dym+1 (m + 1) f (x)) P (xjdym f m +1 a)P (dymjf m + 1 a):
Next note that since x = a(dxm dym), it does not depend directly on f . Consequently we
expand in dxm to remove the f dependence in P (xjdym f m +1 a):
X
X
P (dym+1 jf m +1 a) =
(dym+1(m + 1) f (x)) P (xjdm a) P (dxmjdym f m + 1 a)
f x dxm
f
=
P (dymjf m + 1 a)
X
f dxm
(dym+1(m + 1) f (a(dm))) P (dm jf m a)
23
where use was made of the fact that P (xjdm a) = (x a(dm)) and the fact that P (dm jf m +
1 a) = P (dm jf m a).
The sum over cost functions f is done rst. The cost function is dened both over those
points restricted to dxm and those points outside of dxm . P (dm jf m a) will depend on the f
values dened over points inside dxm while (dym+1 (m + 1) f (a(dm ))) depends only on the f
values dened over points outside dxm . (Recall that a(dxm) 62 dxm .) So we have
X
X X
X y
P (dym+1 jf m +1 a) =
P (dm jf m a)
(dm+1 (m +1) f (a(dm))): (8)
dxm f (x2dxm )
f (x62dxm)
The sum Pf (x62dxm) contributes a constant, jYjjXj;m;1 , equal to the number
dened over points not in dxm passing through (dxm+1 (m + 1) f (a(dm ))). So
X
X
P (dym+1 jf m +1 a) = jYjjXj;m;1
P (dm jf m a)
f
f (x2dxm ) dxm
1 X P (d jf m a)
= jYj
m
f dxm
1 X P (dy jf m a)
= jYj
m
f
f
of functions
By hypothesis the right hand side of this equation is independent of a, so the left hand side
must also be. This completes the proof.
B NFL proof for time-dependent cost functions
In analogy with the proof of the static NFL theorem, thePproof for the time-dependent case
proceeds by establishing the a-independence of the sum T P (cj f T m a), where here c is
either dym or Dmy .
To begin, replace each T in this sum with a set of cost functions, fi, one for each iteration
of the algorithm. To do this, we start with the following:
X
XX X
P (cjf T m a) =
P (cjf~ dxm T m a)P (f2 fm dxm j f1 T m a)
T
=
T dxm f2 fm
X X
dxm f2 fm
X
P (~c j f~ dxm )P (dxm j f~ m a) P (f2 fm j f1 T m a)
T
where the sequence of cost functions, fi, has been indicated by the vector f~ = (f1 fm).
In the next step, the sum over all possible T is decomposed into a series of sums. Each sum
in the series is over the values T can take for one particular iteration of the algorithm. More
formally, using fi+1 = Ti(fi), we write
X
X X
P (cjf T m a) =
P (~c j f~ dxm )P (dxm j f~ m a)
T
dxm f2 fm
X
(f2 T1(f1)) T1
X
Tm
24
1
;
(fm Tm;1(Tm;2( T1(f1)))):
Note that PT P (cjf T m a) is independent of the values of Ti>m;1, so those values can be
absorbed into an overall a-independent proportionality constant.
Consider the innermost sum over Tm;1, for xed values of the outer sum indices T1 : : :Tm;2.
For xed values of the outer indices Tm;1(Tm;2( T1(f1))) is just a particular xed cost function. Accordingly, the innermost sum over Tm;1 is simply the number of bijections of F that
map that xed cost function to fm . This is the constant, (jFj; 1)!. Consequently, evaluating
the Tm;1 sum yields
X
X X
P (cj f T m a1) /
P (cjf~ dxm)P (dxm j f~ m a)
dxm f2 fm
X
T
(f2 T1(f1)) T1
X
Tm
2
(fm;1 Tm;2(Tm;3( T1(f1)))):
;
The sum over Tm;2 can be accomplished in the same manner Tm;1 is summed over. In fact,
all the sums over all Ti can be done, leaving
X
X X
P (cjf T m a1) /
P (Dmy jf~ dxm )P (dxm j f~ m a)
T
=
dxm f2 fm
X X
P (cjf~ dxm)P (dxm j f1 fm;1 m a):
dxm f2 fm
(9)
In this last step the statistical independence of c and fm has been used.
Further progress depends on whether c represents dym or Dmy . We begin with analysis of
the Dmy case. For this case P (cjf~ dxm) = P (Dmy jfm dxm ), since Dmy only reects cost values
from the last cost function, fm. Using this result gives
X
X X
X
P (Dmy j f T m a1) /
P (dxm jf1 fm;1 m a) P (Dmy jfm dxm)
dxm f2 fm
T
fm
1
;
The nal sum over fm is a constant equal to the number of ways of generating the sample
Dmy from cost values drawn from fm . The important point is that it is independent of
the particular dxm . Because of this the sum over dxm can be evaluated eliminating the a
dependence.
X
X X
P (Dmy jf T m a) /
P (dxm j f1 fm;1 m a) / 1
T
f2 fm
;
1
dxm
This completes the proof of Theorem 2 for the case of Dmy .
The proof of Theorem 2 is completed by turning to the dym case. This is considerably
more dicult since P (~c j f~ dxm ) can not be simplied so that the sums over fi can not be
decoupled. Nevertheless, the NFL result still holds. This is proven by expanding Equation
(9) over possible dym values.
X X X
X
P (dym jdym)P (dym j f~ dxm)P (dxm j f1 fm;1 m a)
P (dym jf T m a) /
T
dxm f2 fm dym
X X
X
P (dxm
= P (dymjdym )
dxm f2 fm
dym
25
m
Y
j f1 fm;1 m a) (dym(i) fi(dxm (i)))
i=1
(10)
The
innermost sum over fm only has an eect on the (dym(i) fi(dxm(i))) term so it contributes
P
y
x
jXj;1 . This leaves
fm (dm (m) fm(dm (m))). This is a constant, equal to jYj
X
T
P (dym j f T m a) /
X
dym
P (dym jdym)
X
dxm f2 fm
The sum over dxm(m) is now simple,
X
T
P (dymjf T m a) /
X
dym
P (dym jdym)
mY
;1
i=1
X
1
P (dxm j f1 fm;1 m a)
;
X
X
X
dxm (1) dxm (m;1) f2fm
;
1
mY
;1
i=1
(dym(i) fi(dxm(i))):
P (dxm;1 j f1 fm;2 m a)
(dym(i) fi(dxm (i))):
The above equation is of the same form as Equation (10), only with a remaining population
of size m ; 1 rather than m. Consequently, in an analogous manner to the scheme used to
evaluate the sums over fm and dxm (m) that existed in Equation (10), the sums over fm;1 and
dxm (m ; 1) can be evaluated. Doing so simply generates more a-independent proportionality
constants. Continuing in this manner, all sums over the fi can be evaluated to nd
X
T
P (~c j f T m a1) /
X
dym
P (~c j dym )
X
dxm (1)
P (dxm (1) j m a) (dym(1) f1(dxm (1))):
There is algorithm-dependence in this result but it is the trivial dependence discussed previously. It arises from how the algorithm selects the rst x point in its population, dxm (1).
Restricting interest to those points in the sample that are generated subsequent to the rst,
this result shows that there are no distinctions between algorithms. Alternatively, summing
over the initial cost function f1, all points in the sample could be considered while still
retaining an NFL result.
C Proof of
f result
As noted in the discussion leading up to Theorem 3 the fraction of functions giving a specied
histogram ~c = m~ is independent of the algorithm. Consequently, a simple algorithm is
used to prove the theorem. The algorithm visits points in X in some canonical order,
say x1 x2 : : : xm. Recall that the histogram ~c is specied by giving the frequencies of
occurrence, across the x1 x2 : : : xm, for each of the jYj possible cost values. The number
of f 's giving the desired histogram under this algorithm is just the multinomial giving the
number of ways of distributing the cost values in ~c. At the remaining jXj ; m points in X
the cost can assume any of the jYj f values giving the rst result of Theorem 3.
The expression of f (~ ) in terms of the entropy of ~ follows from an application of
Stirling's approximation to order O(1=m), which is valid when all of the ci are large. In this
26
case the multinomial is written:
m
!
jYj
jYj
X
X
1
ln c c c = m ln m ; ci ln ci + 2 ln m ; ln ci
1 2
jYj
i=1
i=1
jYj
= m S (~ ) + 1 1 ; jYj ln m ; X ln i
2
i=1
from which the theorem follows by exponentiating this result.
D Proof of
alg result
In this section the proportion of all algorithms that give a particular ~c for a particular f is
calculated. The calculation proceeds in several steps:
Since X is nite there are nite number of dierent samples. Therefore any (deterministic) a is a huge { but nite { list indexed by all possible d's. Each entry in the list is the x
the a in question outputs for that d-index.
Consider any particular unordered set of m (X Y ) pairs where no two of the pairs share
the same x value. Such a set is called an unordered path . Without loss of generality, from
now on we implicitly restrict the discussion to unordered paths of length m. A particular
is in or from a particular f if there is a unordered set of m (x f (x)) pairs identical to .
The numerator on the right-hand side of Equation (3) is the number of unordered paths in
the given f that give the desired ~c.
The number of unordered paths in f that give the desired ~c - the numerator on the
right-hand side of Equation (3) - is proportional to the number of a's that give the desired
~c for f and the proof of this claim constitutes a proof of Equation (3).) Furthermore, the
proportionality constant is independent of f and ~c.
Proof: The proof is established by constructing a mapping : a 7! taking in an a that
gives the desired ~c for f , and producing a that is in f and gives the desired ~c. Showing
that for any the number of algorithms a such that (a) = is a constant, independent of
f , and ~c. and that is single-valued will complete the proof.
Recalling that that every x value in an unordered path is distinct any unordered path gives a set of m! dierent ordered paths. Each such ordered path ord in turn provides a set
of m successive d's (if the empty d is included) and a following x. Indicate by d(ord) this
set of the rst m d's provided by ord.
>From any ordered path ord a \partial algorithm" can be constructed. This consists
of the list of an a, but with only the m d(ord ) entries in the list lled in, the remaining
entries are blank. Since there are m! distinct partial a's for each (one for each ordered path
corresponding to ), there are m! such partially lled-in lists for each . A partial algorithm
may or may not be consistent with a particular full algorithm. This allows the denition
of the inverse of : for any that is in f and gives ~c, ;1() (the set of all a that are
consistent with at least one partial algorithm generated from and that give ~c when run on
f ).
27
To complete the rst part of the proof it must be shown that for all that are in f and
give ~c, ;1() contains the same number of elements, regardless of , f , or c. To that end,
rst generate all ordered paths induced by and then associate each such ordered path with
a distinct m-element partial algorithm. Now how many full algorithms lists are consistent
with at least one of these partial algorithm partial lists? How this question is answered is
the core of this appendix. To answer this question, reorder the entries in each of the partial
algorithm lists by permuting the indices d of all the lists. Obviously such a reordering won't
change the answer to our question.
Reordering is accomplished by interchanging pairs of d indices. First, interchange any
d index of the form ((dxm(1) dym (1)) : : : (dxm (i m) dym(i m))) whose entry is lled in
in any of our partial algorithm lists with d0(d) ((dxm(1) z) : : : (dxm (i) z)), where z is
some arbitrary constant Y value and xj refers to the j 'th element of X . Next, create some
arbitrary but xed ordering of all x 2 X : (x1 : : : xjXj). Then interchange any d0 index of
the form ((dxm(1) z : : : (dxm (i m) z) whose entry is lled in in any of our (new) partial
algorithm lists with d00(d0) ((x1 z) : : : (xm z)). Recall that all the dxm(i) must be distinct.
By construction, the resultant partial algorithm lists are independent of , ~c and f , as is the
number of such lists (it's m!). Therefore the number of algorithms consistent with at least
one partial algorithm list in ;1() is independent of , c and f . This completes the rst
part of the proof.
For the second part, rst choose any 2 unordered paths that dier from one another, A
and B . There is no ordered path Aord constructed from A that equals an ordered path Bord
constructed from B . So choose any such Aord and any such Bord. If they disagree for the
null d, then we know that there is no (deterministic) a that agrees with both of them. If
they agree for the null d, then since they are sampled from the same f , they have the same
single-element d. If they disagree for that d, then there is no a that agrees with both of
them. If they agree for that d, then they have the same double-element d. Continue in this
manner all the up to the (m ; 1)-element d. Since the two ordered paths dier, they must
have disagreed at some point by now, and therefore there is no a that agrees with both of
them. Since this is true for any Aord from A and any Bord from B , we see that there is no a
in ;1(A) that is also in ;1(B ). This completes the proof.
To show the relation to the Kullback-Liebler distance the product of binomials is expanded with the aid of Stirlings approximation when both Ni and ci are large:
jYj
Y
!
jYj
Y
!
jYj
N
i X 1
ln
= ; 2 ln 2 + Ni ln Ni ; ci ln ci ; (Ni ; ci) ln(Ni ; ci) +
i=1 ci
i=1
1 ln N ; ln(N ; c ) ; ln c :
i
i
i
i
2
We it has been assumed that ci=Ni 1, which is reasonable when m jXj. Expanding
ln(1 ; z) = ;z ; z2=2 ; , to second order gives
ln
i=1
jYj
Ni X
Ni ; 1 ln c + c ; 1 ln 2 ; ci c ; 1 + c
ln
=
i
ci 2 i i 2
2Ni i
ci
i=1
28
Using m=jXj 1 then in terms of ~ and ~ one nds
ln
jYj
Y
i=1
!
Ni ;mD (~ ~ ) + m ; m ln m ; jYj ln 2
KL
ci =
jXj 2
jYj
X
m i (1 ; m + )
; 21 ln(im) + 2jXj
i
i
i=1
where DKL (~ ~ ) Pi i ln(i=i) is the Kullback-Liebler distance between the distributions
~ and ~ . Exponentiating this expression yields the second result in Theorem 4.
E Benchmark measures of performance
The result for each benchmark
measure is established in turn.
P
The rst measure is f P (min(dym )jf m a). Consider
X
f
P (min(dym )jf m a)
(11)
for which the summand equals 0 or 1 for all f and deterministic a. It is 1 only if
i) f (dxm (1)) = dym(1)
ii) f (adm(1)]) = dym (2)
iii) f (adm(1) dm (2)]) = dym (3)
and so on. These restrictions will x the value of f (x) at m points while f remains free at
all other points. Therefore
X
f
P (dym j f m a) = jYjjXj;m:
Using this result in Equation (11) we nd
X
X
P (min(dy ) > j f m) = 1
P (min(dy ) > j dy ) = 1
f
m
jYjm dym
= jYj1 m (jYj ; )m:
m
m
X
jYjm dym3min(dym )> 1
which is the result quoted in Theorem P
5.
m
m
In the limit as jYj gets large write f E (min(~c)jf m) = PjYj
=1 ! ( ; 1) ; ! ()] and
PjYj;1
substitute in for !() = 1 ; =jYj. Replacing with + 1 turns the sum into =0 +
+1 )m ]. Next, write jYj = b=& for some b and multiply and divide the
1] (1 ; !Yj )m ; (1 ; jYj
summand by &. Since jYj ! 1 then & ! 0. To take the limit of & ! 0, apply L'hopital's
29
rule to the ratio in the summand. Next use the fact that & is going to 0 to cancel terms
in the summand. Carrying
through the algebra, and dividing by b=&, we get a Riemann
Rb
m
sum of the form b2 0 dx x(1 ; x=b)m;1. Evaluating the integral gives the second result in
Theorem 5.
The second benchmark concerns the behavior of the random algorithm. Marginalizing
over the Y values of dierent histograms ~c, the performance of a~ is
X
P (min(~c) j f m ~a) = P (min(~c) j ~c) P (~c j f m ~a)
~c
Now P (~c j f m ~a) is the probability of obtaining histogram ~c in m random draws from the
histogram N~ of the function f . This
can be viewed as the denition of a~. This probability
QjYj Ni jXj
has been calculated previously as i=1 ci = m . So
1
P (min(~c) j f m ~a) = jXj
=
m
X
m
X
jYj
X
( ci
c
=0
c
=0
i=1
1
m
jYj
m
m
X
X
1 X
(
ci
jXj
c
=0
c
=0
i
=
m P
()jXj
i= Ni
m
m
jXj jXj
m
m
m)P (min(~c) j~c)
jY j
jY j
m)
jYj N !
Y
i
i=
jYj
Y
i=1
Ni
ci
!
ci
jY j
=
which is Equation (4) of Theorem 6.
F Proof related to minimax distinctions between algorithms
The proof is by example.
Consider three points in X , x1 x2, and x3, and three points in Y , y1 y2, and y3.
1) Let the rst point a1 visits be x1, and the rst point a2 visits be x2.
2) If at its rst point a1 sees a y1 or a y2, it jumps to x2. Otherwise it jumps to x3.
3) If at its rst point a2 sees a y1, it jumps to x1. If it sees a y2, it jumps to x3.
Consider the cost function that has as the Y values for the three X values fy1 y2 y3g,
respectively.
For m = 2, a1 will produce a population (y1 y2) for this function, and a2 will produce
(y2 y3).
The proof is completed if we show that there is no cost function so that a1 produces a
population containing y2 and y3 and such that a2 produces a population containing y1 and
y2.
There are four possible pairs of populations to consider:
30
i) (y2 y3) (y1 y2)]
ii) (y2 y3) (y2 y1)]
iii) (y3 y2) (y1 y2)]
iv) (y3 y2) (y2 y1)].
Since if its rst point is a y2 a1 jumps to x2 which is where a2 starts, when a1's rst point is
a y2 its second point must equal a2's rst point. This rules out possibilities i) and ii).
For possibilities iii) and iv), by a1's population we know that f must be of the form
fy3 s y2g, for some variable s. For case iii), s would need to equal y1, due to the rst point
in a2's population. However for that case, the second point a2 sees would be the value at x1,
which is y3, contrary to hypothesis.
For case iv), we know that the s would have to equal y2, due to the rst point in a2's
population. However that would mean that a2 jumps to x3 for its second point, and would
therefore see a y2, contrary to hypothesis.
Accordingly, none of the four cases is possible. This is a case both where there is no
symmetry under exchange of dy 's between a1 and a2, and no symmetry under exchange of
histograms. QED.
G Fixed cost functions and choosing procedures
Since any deterministic search algorithm is a mapping from d D to x X , any search
algorithm is a vector in the space X D . The components of such a vector are indexed by the
possible populations, and the value for each component is the x that the algorithm produces
given the associated population.
Consider now a particular population d of size m. Given d, we can say whether any
other population of size greater than m has the (ordered) elements of d as its rst m (ordered) elements. The set of those populations that do start with d this way denes a set of
components of any algorithm vector a. Those components will be indicated by ad.
The remaining components of a are of two types. The rst is given by those populations
that are equivalent to the rst M < m elements in d for some M . The values of those
components for the vector algorithm a will be indicated by ad. The second type consists of
those components corresponding to all remaining populations. Intuitively, these are populations that are not compatible with d. Some examples of such populations are populations
that contain as one of their rst m elements an element not found in d, and populations that
re-order the elements found in d. The values of a for components of this second type will be
indicated by a?d.
Let proc be either A or B . We are interested in
X
aa
0
P (c>m j f d1 d2 k a a0 proc) =
X
X
X
a da d a da d a da d
?
0
? 0
31
0
0
0
0
P (c>m jf d d0 k a a0 proc):
The summand is independent of the values of a?d and a0?d for either of our two d's. In
addition, the number of such values is a constant. (It is given by the product, over all
populations not consistent with d, of the number of possible x each such population could
be mapped to.) Therefore, up to an overall constant independent of d, d0, f , and proc, the
sum equals
X
X
a da d a da d
0
0
0
0
P (c>m j f d d0 ad a0d ad a0d proc):
0
0
By denition, we are implicitly restricting the sum to those a and a0 so that our summand
is dened. This means that we actually only allow one value for each component in ad
(namely, the value that gives the next x element in d), and similarly for a0d . Therefore the
sum reduces to
0
X
a da d
0
P (c>m j f d d0 ad a0d proc):
0
0
Note that no component of ad lies in dx. The same is true of a0d . So the sum over ad is
over the same components of a as the sum over a0d is of a0. Now for xed d and d0, proc's
choice of a or a0 is xed. Accordingly, without loss of generality, the sum can be rewritten
as
0
0
X
a d
P (c>m j f d d0 ad)
with the implicit assumption that c>m is set by ad . This sum is independent of proc.
H Proof of Theorem 9
Let proc refer to a choosing procedure. We are interested in
X
aa
0
P (c>m j f m k a a0 proc) =
X
aa dd
0
0
P (c>m j f d d0 k a a0 proc)
P (d d0 j f k m a a0 proc):
The sum over d and d0 can be moved outside the sum over a and a0. Consider any term in that
sum (i.e., any particular pair of values of d and d0). For that term, P (d d0 j f k m a a0 proc)
is just 1 for those a and a0 that result in d and d0 respectively when run on f , and 0
otherwise. (Recall the assumption that a and a0 are deterministic.) This means that the
P (d d0 j f k m a a0 proc) factor simply restricts our sum over a and a0 to the a and a0
considered in our theorem. Accordingly, our theorem tell us that the summand of the sum
over d and d0 is the same for choosing procedures A and B . Therefore the full sum is the
same for both procedures.
32
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