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Given a set of variables, x's. I want to find the values of coefficients for this equation:
y = a_1*x_1 +... +a_n*x_n + c
where a_1,a_2,...,a_n are all unknowns. Thinking this in perspective of data frame, I want to create this value of y for every rows in the data.
My question is: for y, a_1...a_n and c are all unknown, is there a way for me to find a set of solutions a_1,...,a_n under the condition that corr(y,x_1), corr(y,x_2) .... corr(y,x_n) are all greater than 0.7. For simplicity take correlation here as Pearson correlation. I know there would no be unique solution. But how can I construct a set of solutions for a_1,...,a_n to fulfill this condition?
Spent a day to search the idea but could not get any information out of it. Any programming language to tackle this problem is welcomed or at least some reference for this.
No, it is not possible in general. It may be possible in some special cases.
Given x₁, x₂, ... you want to find y = a₁x₁ + a₂x₂ + ... + c so that all the correlations between y and the x's are greater than some target R. Since the correlation is
Corr(y, xi) = Cov(y, xi) / Sqrt[ Var(y) * Var(xi) ]
your constraint is
Cov(y, xi) / Sqrt[ Var(y) * Var(xi) ] > R
which can be rearranged to
Cov(y, xi)² > R² * Var(y) * Var(xi)
and this needs to be true for all i.
Consider the simple case where there are only two columns x₁ and x₂, and further assume that they both have mean zero (so you can ignore the constant c) and variance 1, and that they are uncorrelated. In that case y = a₁x₁ + a₂x₂ and the covariances and variances are
Cov(y, x₁) = a₁
Cov(y, x₂) = a₂
Var(x₁) = 1
Var(x₂) = 1
Var(y) = (a₁)² + (a₂)²
so you need to simultaneously satisfy
(a₁)² > R² * ((a₁)² + (a₂)²)
(a₂)² > R² * ((a₁)² + (a₂)²)
Adding these inequalities together, you get
(a₁)² + (a₂)² > 2 * R² * ((a₁)² + (a₂)²)
which means that in order to satisfy both of the inequalities, you must have R < Sqrt(1/2) (by cancelling common factors on both sides of the inequality). So the very best you could do in this simple case is to choose a₁ = a₂ (the exact value doesn't matter as long as they are equal) and both of the correlations Corr(y,a₁) and Corr(y,a₂) will be equal to 0.707. You cannot achieve correlations higher than this between y and all of the x's simultaneously in this case.
For the more general case with n columns (each of which has mean zero, variance 1 and zero correlation between columns) you cannot simultaneously achieve correlations greater than 1 / sqrt(n) (as pointed out in the comments by #kazemakase).
In general, the more independent variables there are, the lower the correlation you will be able to achieve between y and the x's. Also (although I haven't mentioned it above) the correlations between the x's matter. If they are in general positively correlated, you will be able to achieve a higher target correlation between y and the x's. If they are in general uncorrelated or negatively correlated, you will only be able to achieve low correlations between y and the x's.
I am not expert in this field so read with extreme prejudice!
I am a bit confused by your y
Your y is a single constant and you want to have the correlation between it and all the x_i values be > 0.7 ? I am no math/statistics expert but my feelings for this are that this is achievable only if the correlation between x_i,x_j upholds the same condition. in that case you can simply do the average of x_i like this:
y=(x_1+x_2+x_3+...+x_n)/n
so the a_i=1.0/n and c=0.0 But still the question is:
What meaning has a correlation between 2 numbers only?
More reasonable would be if y is a function dependent on x
for example like this:
y(x) = a_1*(x-x_1)+... +a_n*(x-x_n) + c
or any other equation (hard to make any without knowing where it came from and for what purpose). Then you can compute the correlation between two sets
X = { x_1 , x_2 ,..., x_n }
Y = { y(x_1),y(x_2),...y(x_n) }
In that case I would give try approximation search for the c,a_i constants to maximize correlation between X,Y, but the results complexity for the whole thing would be insane. So instead I would tweak just one constant. at the time
set some safe c,a_1,a_2,... constants
tweak a_1
so compute correlation for (a_1-delta) and (a_1+delta) and then choose the direction which is in favor of correlation. then keep going in that direction until the correlation coefficient start to drop.
Then you can recursively to this again with smaller delta. Btw this is exactly what my approx class does from the link above.
loop #2 through all the a_i
loop this whole few times to enhance precision
May be you could compute the c after each run to minimize the distance between X,Y sets.
I have number of strings (n strings) and I am computing edit distance between strings in a way that I take first one and compare it to the (n-1) remaining strings, second one and compare it to (n-2) remaining, ..., comparing until I ran out of the strings.
Why would an average edit distance be computed as sum of all the edit distances between all the strings divided by the number of comparisons squared. This squaring is confusing me.
Thanks,
Jannine
I assume you have somewhere an answer that seems to come with a squared factor -which I'll take as n^2, where n is the number of strings (not the number of distinct comparisons, which is n*(n-1)/2, as +flaschenpost points to ). It would be easier to give you a more precise answer if you'd exactly quote what that answer is.
From what I understand of your question, it isn't, at least it's not the usual sample average. It is, however, a valid estimator of central tendency with the caveat that it is a biased estimator.
See https://en.wikipedia.org/wiki/Bias_of_an_estimator.
Let's define the sample average, which I will denote as X', by
X' = \sum^m_i X_i/N
IF N=m, we get the standard average. In your case, this is the number of distinct pairs which is m=n*(n-1)/2. Let's call this average Xo.
Then if N=n*n, it is
X' = (n-1)/(2*n) Xo
Xo is an unbiased estimator of the population mean \mu. Therefore, X' is biased by a factor f=(n-1)/(2*n). For n very large this bias tends to 1/2.
That said, it could be that the answer you see has a sum that runs not just over distinct pairs. The normalization would then change, of course. For instance, we could extend that sum to all pairs without changing the average value: The correct normalization would then be N = n*(n-1); the value of the average would still be Xo though as the number of summands has double as well.
Those things are getting easier to understand if done by hand with pen and paper for a small example.
If you have the 7 Strings named a,b,c,d,e,f,g, then the simplest version would
Compare a to b, a to c, ... , a to g (this are 6)
Compare b to a, b to c, ... , b to g (this are 6)
. . .
Compare g to a, g to b, ... , g to f (this are 6)
So you have 7*6 or n*(n-1) values, so you divide by nearly 7^2. This is where the square comes from. Maybe you even compare a to a, which should bring a distance of 0 and increase the values to 7*7 or n*n. But I would count it a bit as cheating for the average distance.
You could double the speed of the algorithm, just changing it a small bit
Compare a to b, a to c, ... , a to g (this are 6)
Compare b to c, ... , b to g (this are 5)
Compare c to d, ... , b to g (this are 4)
. . .
Compare f to g (this is 1)
That is following good ol' Gauss 7*6/2, or n*(n-1)/2.
So in Essence: Try doing a simple example on paper and then count your distance values.
Since Average is still and very simply the same as ever:
sum(values) / count(values)
I have polynomials of nontrivial degree (4+) and need to robustly and efficiently determine whether or not they have a root in the interval [0,T]. The precise location or number of roots don't concern me, I just need to know if there is at least one.
Right now I'm using interval arithmetic as a quick check to see if I can prove that no roots can exist. If I can't, I'm using Jenkins-Traub to solve for all of the polynomial roots. This is obviously inefficient since it's checking for all real roots and finding their exact positions, information I don't end up needing.
Is there a standard algorithm I should be using? If not, are there any other efficient checks I could do before doing a full Jenkins-Traub solve for all roots?
For example, one optimization I could do is to check if my polynomial f(t) has the same sign at 0 and T. If not, there is obviously a root in the interval. If so, I can solve for the roots of f'(t) and evaluate f at all roots of f' in the interval [0,T]. f(t) has no root in that interval if and only if all of these evaluations have the same sign as f(0) and f(T). This reduces the degree of the polynomial I have to root-find by one. Not a huge optimization, but perhaps better than nothing.
Sturm's theorem lets you calculate the number of real roots in the range (a, b). Given the number of roots, you know if there is at least one. From the bottom half of page 4 of this paper:
Let f(x) be a real polynomial. Denote it by f0(x) and its derivative f′(x) by f1(x). Proceed as in Euclid's algorithm to find
f0(x) = q1(x) · f1(x) − f2(x),
f1(x) = q2(x) · f2(x) − f3(x),
.
.
.
fk−2(x) = qk−1(x) · fk−1(x) − fk,
where fk is a constant, and for 1 ≤ i ≤ k, fi(x) is of degree lower than that of fi−1(x). The signs of the remainders are negated from those in the Euclid algorithm.
Note that the last non-vanishing remainder fk (or fk−1 when fk = 0) is a greatest common
divisor of f(x) and f′(x). The sequence f0, f1,. . ., fk (or fk−1 when fk = 0) is called a Sturm sequence for the polynomial f.
Theorem 1 (Sturm's Theorem) The number of distinct real zeros of a polynomial f(x) with
real coefficients in (a, b) is equal to the excess of the number of changes of sign in the sequence f0(a), ..., fk−1(a), fk over the number of changes of sign in the sequence f0(b), ..., fk−1(b), fk.
You could certainly do binary search on your interval arithmetic. Start with [0,T] and substitute it into your polynomial. If the result interval does not contain 0, you're done. If it does, divide the interval in 2 and recurse on each half. This scheme will find the approximate location of each root pretty quickly.
If you eventually get 4 separate intervals with a root, you know you are done. Otherwise, I think you need to get to intervals [x,y] where f'([x,y]) does not contain zero, meaning that the function is monotonically increasing or decreasing and hence contains at most one zero. Double roots might present a problem, I'd have to think more about that.
Edit: if you suspect a multiple root, find roots of f' using the same procedure.
Use Descartes rule of signs to glean some information. Just count the number of sign changes in the coefficients. This gives you an upper bound on the number of positive real roots. Consider the polynomial P.
P = 131.1 - 73.1*x + 52.425*x^2 - 62.875*x^3 - 69.225*x^4 + 11.225*x^5 + 9.45*x^6 + x^7
In fact, I've constructed P to have a simple list of roots. They are...
{-6, -4.75, -2, 1, 2.3, -i, +i}
Can we determine if there is a root in the interval [0,3]? Note that there is no sign change in the value of P at the endpoints.
P(0) = 131.1
P(3) = 4882.5
How many sign changes are there in the coefficients of P? There are 4 sign changes, so there may be as many as 4 positive roots.
But, now substitute x+3 for x into P. Thus
Q(x) = P(x+3) = ...
4882.5 + 14494.75*x + 15363.9*x^2 + 8054.675*x^3 + 2319.9*x^4 + 370.325*x^5 + 30.45*x^6 + x^7
See that Q(x) has NO sign changes in the coefficients. All of the coefficients are positive values. Therefore there can be no roots larger than 3.
So there MAY be either 2 or 4 roots in the interval [0,3].
At least this tells you whether to bother looking at all. Of course, if the function has opposite signs on each end of the interval, we know there are an odd number of roots in that interval.
It's not that efficient, but is quite reliable. You can construct the polynomial's Companion Matrix (A sparse matrix whose eigenvalues are the polynomial's roots).
There are efficient eigenvalue algorithms that can find eigenvalues in a given interval. One of them is the inverse iteration (Can find eigenvalues closest to some input value. Just give the middle point of the interval as the above value).
If the value f(0)*f(t)<=0 then you are guaranteed to have a root. Otherwise you can start splitting the domain into two parts (bisection) and check the values in the ends until you are confident there is no root in that segment.
if f(0)*f(t)>0 you either have no, two, four, .. roots. Your limit is the polynomial order. if f(0)*f(t)<0 you may have one, three, five, .. roots.
suppose I have the following 2 random variables :
X where mean = 6 and stdev = 3.5
Y where mean = -42 and stdev = 5
I would like to create a new random variable Z based on the first two and knowing that : X happens 90% of the time and Y happens 10% of the time.
It is easy to calculate the mean for Z : 0.9 * 6 + 0.1 * -42 = 1.2
But is it possible to generate random values for Z in a single function?
Of course, I could do something along those lines :
if (randIntBetween(1,10) > 1)
GenerateRandomNormalValue(6, 3.5);
else
GenerateRandomNormalValue(-42, 5);
But I would really like to have a single function that would act as a probability density function for such a random variable (Z) that is not necessary normal.
sorry for the crappy pseudo-code
Thanks for your help!
Edit : here would be one concrete interrogation :
Let's say we add the result of 5 consecutives values from Z. What would be the probability of ending with a number higher than 10?
But I would really like to have a
single function that would act as a
probability density function for such
a random variable (Z) that is not
necessary normal.
Okay, if you want the density, here it is:
rho = 0.9 * density_of_x + 0.1 * density_of_y
But you cannot sample from this density if you don't 1) compute its CDF (cumbersome, but not infeasible) 2) invert it (you will need a numerical solver for this). Or you can do rejection sampling (or variants, eg. importance sampling). This is costly, and cumbersome to get right.
So you should go for the "if" statement (ie. call the generator 3 times), except if you have a very strong reason not to (using quasi-random sequences for instance).
If a random variable is denoted x=(mean,stdev) then the following algebra applies
number * x = ( number*mean, number*stdev )
x1 + x2 = ( mean1+mean2, sqrt(stdev1^2+stdev2^2) )
so for the case of X = (mx,sx), Y= (my,sy) the linear combination is
Z = w1*X + w2*Y = (w1*mx,w1*sx) + (w2*my,w2*sy) =
( w1*mx+w2*my, sqrt( (w1*sx)^2+(w2*sy)^2 ) ) =
( 1.2, 3.19 )
link: Normal Distribution look for Miscellaneous section, item 1.
PS. Sorry for the wierd notation. The new standard deviation is calculated by something similar to the pythagorian theorem. It is the square root of the sum of squares.
This is the form of the distribution:
ListPlot[BinCounts[Table[If[RandomReal[] < .9,
RandomReal[NormalDistribution[6, 3.5]],
RandomReal[NormalDistribution[-42, 5]]], {1000000}], {-60, 20, .1}],
PlotRange -> Full, DataRange -> {-60, 20}]
It is NOT Normal, as you are not adding Normal variables, but just choosing one or the other with certain probability.
Edit
This is the curve for adding five vars with this distribution:
The upper and lower peaks represent taking one of the distributions alone, and the middle peak accounts for the mixing.
The most straightforward and generically applicable solution is to simulate the problem:
Run the piecewise function you have 1,000,000 (just a high number) of times, generate a histogram of the results (by splitting them into bins, and divide the count for each bin by your N (1,000,000 in my example). This will leave you with an approximation for the PDF of Z at every given bin.
Lots of unknowns here, but essentially you just wish to add the two (or more) probability functions to one another.
For any given probability function you could calculate a random number with that density by calculating the area under the probability curve (the integral) and then generating a random number between 0 and that area. Then move along the curve until the area is equal to your random number and use that as your value.
This process can then be generalized to any function (or sum of two or more functions).
Elaboration:
If you have a distribution function f(x) which ranges from 0 to 1. You could calculate a random number based on the distribution by calculating the integral of f(x) from 0 to 1, giving you the area under the curve, lets call it A.
Now, you generate a random number between 0 and A, let's call that number, r. Now you need to find a value t, such that the integral of f(x) from 0 to t is equal to r. t is your random number.
This process can be used for any probability density function f(x). Including the sum of two (or more) probability density functions.
I'm not sure what your functions look like, so not sure if you are able to calculate analytic solutions for all this, but worse case scenario, you could use numeric techniques to approximate the effect.
This question on getting random values from a finite set got me thinking...
It's fairly common for people to want to retrieve X unique values from a set of Y values. For example, I may want to deal a hand from a deck of cards. I want 5 cards, and I want them to all be unique.
Now, I can do this naively, by picking a random card 5 times, and try again each time I get a duplicate, until I get 5 cards. This isn't so great, however, for large numbers of values from large sets. If I wanted 999,999 values from a set of 1,000,000, for instance, this method gets very bad.
The question is: how bad? I'm looking for someone to explain an O() value. Getting the xth number will take y attempts...but how many? I know how to figure this out for any given value, but is there a straightforward way to generalize this for the whole series and get an O() value?
(The question is not: "how can I improve this?" because it's relatively easy to fix, and I'm sure it's been covered many times elsewhere.)
Variables
n = the total amount of items in the set
m = the amount of unique values that are to be retrieved from the set of n items
d(i) = the expected amount of tries needed to achieve a value in step i
i = denotes one specific step. i ∈ [0, n-1]
T(m,n) = expected total amount of tries for selecting m unique items from a set of n items using the naive algorithm
Reasoning
The first step, i=0, is trivial. No matter which value we choose, we get a unique one at the first attempt. Hence:
d(0) = 1
In the second step, i=1, we at least need 1 try (the try where we pick a valid unique value). On top of this, there is a chance that we choose the wrong value. This chance is (amount of previously picked items)/(total amount of items). In this case 1/n. In the case where we picked the wrong item, there is a 1/n chance we may pick the wrong item again. Multiplying this by 1/n, since that is the combined probability that we pick wrong both times, gives (1/n)2. To understand this, it is helpful to draw a decision tree. Having picked a non-unique item twice, there is a probability that we will do it again. This results in the addition of (1/n)3 to the total expected amounts of tries in step i=1. Each time we pick the wrong number, there is a chance we might pick the wrong number again. This results in:
d(1) = 1 + 1/n + (1/n)2 + (1/n)3 + (1/n)4 + ...
Similarly, in the general i:th step, the chance to pick the wrong item in one choice is i/n, resulting in:
d(i) = 1 + i/n + (i/n)2 + (i/n)3 + (i/n)4 + ... = = sum( (i/n)k ), where k ∈ [0,∞]
This is a geometric sequence and hence it is easy to compute it's sum:
d(i) = (1 - i/n)-1
The overall complexity is then computed by summing the expected amount of tries in each step:
T(m,n) = sum ( d(i) ), where i ∈ [0,m-1] = = 1 + (1 - 1/n)-1 + (1 - 2/n)-1 + (1 - 3/n)-1 + ... + (1 - (m-1)/n)-1
Extending the fractions in the series above by n, we get:
T(m,n) = n/n + n/(n-1) + n/(n-2) + n/(n-3) + ... + n/(n-m+2) + n/(n-m+1)
We can use the fact that:
n/n ≤ n/(n-1) ≤ n/(n-2) ≤ n/(n-3) ≤ ... ≤ n/(n-m+2) ≤ n/(n-m+1)
Since the series has m terms, and each term satisfies the inequality above, we get:
T(m,n) ≤ n/(n-m+1) + n/(n-m+1) + n/(n-m+1) + n/(n-m+1) + ... + n/(n-m+1) + n/(n-m+1) = = m*n/(n-m+1)
It might be(and probably is) possible to establish a slightly stricter upper bound by using some technique to evaluate the series instead of bounding by the rough method of (amount of terms) * (biggest term)
Conclusion
This would mean that the Big-O order is O(m*n/(n-m+1)). I see no possible way to simplify this expression from the way it is.
Looking back at the result to check if it makes sense, we see that, if n is constant, and m gets closer and closer to n, the results will quickly increase, since the denominator gets very small. This is what we'd expect, if we for example consider the example given in the question about selecting "999,999 values from a set of 1,000,000". If we instead let m be constant and n grow really, really large, the complexity will converge towards O(m) in the limit n → ∞. This is also what we'd expect, since while chosing a constant number of items from a "close to" infinitely sized set the probability of choosing a previously chosen value is basically 0. I.e. We need m tries independently of n since there are no collisions.
If you already have chosen i values then the probability that you pick a new one from a set of y values is
(y-i)/y.
Hence the expected number of trials to get (i+1)-th element is
y/(y-i).
Thus the expected number of trials to choose x unique element is the sum
y/y + y/(y-1) + ... + y/(y-x+1)
This can be expressed using harmonic numbers as
y (Hy - Hy-x).
From the wikipedia page you get the approximation
Hx = ln(x) + gamma + O(1/x)
Hence the number of necessary trials to pick x unique elements from a set of y elements
is
y (ln(y) - ln(y-x)) + O(y/(y-x)).
If you need then you can get a more precise approximation by using a more precise approximation for Hx. In particular, when x is small it is possible to
improve the result a lot.
If you're willing to make the assumption that your random number generator will always find a unique value before cycling back to a previously seen value for a given draw, this algorithm is O(m^2), where m is the number of unique values you are drawing.
So, if you are drawing m values from a set of n values, the 1st value will require you to draw at most 1 to get a unique value. The 2nd requires at most 2 (you see the 1st value, then a unique value), the 3rd 3, ... the mth m. Hence in total you require 1 + 2 + 3 + ... + m = [m*(m+1)]/2 = (m^2 + m)/2 draws. This is O(m^2).
Without this assumption, I'm not sure how you can even guarantee the algorithm will complete. It's quite possible (especially with a pseudo-random number generator which may have a cycle), that you will keep seeing the same values over and over and never get to another unique value.
==EDIT==
For the average case:
On your first draw, you will make exactly 1 draw.
On your 2nd draw, you expect to make 1 (the successful draw) + 1/n (the "partial" draw which represents your chance of drawing a repeat)
On your 3rd draw, you expect to make 1 (the successful draw) + 2/n (the "partial" draw...)
...
On your mth draw, you expect to make 1 + (m-1)/n draws.
Thus, you will make 1 + (1 + 1/n) + (1 + 2/n) + ... + (1 + (m-1)/n) draws altogether in the average case.
This equals the sum from i=0 to (m-1) of [1 + i/n]. Let's denote that sum(1 + i/n, i, 0, m-1).
Then:
sum(1 + i/n, i, 0, m-1) = sum(1, i, 0, m-1) + sum(i/n, i, 0, m-1)
= m + sum(i/n, i, 0, m-1)
= m + (1/n) * sum(i, i, 0, m-1)
= m + (1/n)*[(m-1)*m]/2
= (m^2)/(2n) - (m)/(2n) + m
We drop the low order terms and the constants, and we get that this is O(m^2/n), where m is the number to be drawn and n is the size of the list.
There's a beautiful O(n) algorithm for this. It goes as follows. Say you have n items, from which you want to pick m items. I assume the function rand() yields a random real number between 0 and 1. Here's the algorithm:
items_left=n
items_left_to_pick=m
for j=1,...,n
if rand()<=(items_left_to_pick/items_left)
Pick item j
items_left_to_pick=items_left_to_pick-1
end
items_left=items_left-1
end
It can be proved that this algorithm does indeed pick each subset of m items with equal probability, though the proof is non-obvious. Unfortunately, I don't have a reference handy at the moment.
Edit The advantage of this algorithm is that it takes only O(m) memory (assuming the items are simply integers or can be generated on-the-fly) compared to doing a shuffle, which takes O(n) memory.
Your actual question is actually a lot more interesting than what I answered (and harder). I've never been any good at statistitcs (and it's been a while since I did any), but intuitively, I'd say that the run-time complexity of that algorithm would probably something like an exponential. As long as the number of elements picked is small enough compared to the size of the array the collision-rate will be so small that it will be close to linear time, but at some point the number of collisions will probably grow fast and the run-time will go down the drain.
If you want to prove this, I think you'd have to do something moderately clever with the expected number of collisions in function of the wanted number of elements. It might be possible do to by induction as well, but I think going by that route would require more cleverness than the first alternative.
EDIT: After giving it some thought, here's my attempt:
Given an array of m elements, and looking for n random and different elements. It is then easy to see that when we want to pick the ith element, the odds of picking an element we've already visited are (i-1)/m. This is then the expected number of collisions for that particular pick. For picking n elements, the expected number of collisions will be the sum of the number of expected collisions for each pick. We plug this into Wolfram Alpha (sum (i-1)/m, i=1 to n) and we get the answer (n**2 - n)/2m. The average number of picks for our naive algorithm is then n + (n**2 - n)/2m.
Unless my memory fails me completely (which entirely possible, actually), this gives an average-case run-time O(n**2).
The worst case for this algorithm is clearly when you're choosing the full set of N items. This is equivalent to asking: On average, how many times must I roll an N-sided die before each side has come up at least once?
Answer: N * HN, where HN is the Nth harmonic number,
a value famously approximated by log(N).
This means the algorithm in question is N log N.
As a fun example, if you roll an ordinary 6-sided die until you see one of each number, it will take on average 6 H6 = 14.7 rolls.
Before being able to answer this question in details, lets define the framework. Suppose you have a collection {a1, a2, ..., an} of n distinct objects, and want to pick m distinct objects from this set, such that the probability of a given object aj appearing in the result is equal for all objects.
If you have already picked k items, and radomly pick an item from the full set {a1, a2, ..., an}, the probability that the item has not been picked before is (n-k)/n. This means that the number of samples you have to take before you get a new object is (assuming independence of random sampling) geometric with parameter (n-k)/n. Thus the expected number of samples to obtain one extra item is n/(n-k), which is close to 1 if k is small compared to n.
Concluding, if you need m unique objects, randomly selected, this algorithm gives you
n/n + n/(n-1) + n/(n-2) + n/(n-3) + .... + n/(n-(m-1))
which, as Alderath showed, can be estimated by
m*n / (n-m+1).
You can see a little bit more from this formula:
* The expected number of samples to obtain a new unique element increases as the number of already chosen objects increases (which sounds logical).
* You can expect really long computation times when m is close to n, especially if n is large.
In order to obtain m unique members from the set, use a variant of David Knuth's algorithm for obtaining a random permutation. Here, I'll assume that the n objects are stored in an array.
for i = 1..m
k = randInt(i, n)
exchange(i, k)
end
here, randInt samples an integer from {i, i+1, ... n}, and exchange flips two members of the array. You only need to shuffle m times, so the computation time is O(m), whereas the memory is O(n) (although you can adapt it to only save the entries such that a[i] <> i, which would give you O(m) on both time and memory, but with higher constants).
Most people forget that looking up, if the number has already run, also takes a while.
The number of tries nessesary can, as descriped earlier, be evaluated from:
T(n,m) = n(H(n)-H(n-m)) ⪅ n(ln(n)-ln(n-m))
which goes to n*ln(n) for interesting values of m
However, for each of these 'tries' you will have to do a lookup. This might be a simple O(n) runthrough, or something like a binary tree. This will give you a total performance of n^2*ln(n) or n*ln(n)^2.
For smaller values of m (m < n/2), you can do a very good approximation for T(n,m) using the HA-inequation, yielding the formula:
2*m*n/(2*n-m+1)
As m goes to n, this gives a lower bound of O(n) tries and performance O(n^2) or O(n*ln(n)).
All the results are however far better, that I would ever have expected, which shows that the algorithm might actually be just fine in many non critical cases, where you can accept occasional longer running times (when you are unlucky).