Prove the Security of Schnorr's Signature Scheme
Pointcheval and Stern [PS00] proved that the Schnorr signature is existentially unforgeable under chosen-message attacks (EU-CMA) in the random oracle model assuming that the discrete-logarithm problem$^1$ (DLP)  is hard. 

On a high level, the reduction (from DLP to the EU-CMA-security of Schnorr signature) works as follows. The reduction algorithm $\mathcal{B}$ embeds its DLP challenge $g^\alpha$ into the public key (i.e., sets $y=g^\alpha$) and then  uses the oracle-replay attack to obtain, from the forger $\mathcal{F}$, two different forgeries that share the signing randomness ($r=g^k$). This enables $\mathcal{B}$ to solve for $\alpha$.

For simplicity, let's first focus on a weaker model called the existential forgery under no-message attacks (EU-NMA) and a strong forger that is always successful. We show that a strong forger $\mathcal{F}$ that breaks the Schnorr signature in the EU-NMA model making at most $q$ queries to the random oracle $H$ (i.e., a $(1,q)$-adversary) can be used to break the DLP with probability $1/q$. 

This requires two rounds of simulation: 


Round 1. $\mathcal{B}$ runs $\mathcal{F}$ on $(G,g,g^\alpha)$; the random oracle $H$ is simulated in the standard manner (i.e., lazy sampling plus a table to ensure consistency). At the end of this round $\mathcal{F}$ returns a forgery $(M_0^*,(s_0^*,e_0^*))$ where $e_0^*=H(M_0^*\|r_0^*)$. For simplicity, it is assumed that $\mathcal{F}$ made the random oracle query $H(M_0^*\|r_0^*)$. 
Rounds 2. Now $\mathcal{B}$ rewinds $\mathcal{F}$ to the point where it made the random oracle query  $H(M_0^*\|r_0^*)$ in Round 1 (the "critical" point) and re-runs $\mathcal{F}$ but answering the fresh random oracle queries independent of the previous round.$^2$ This constitutes the oracle-replay attack. At the end of Round 2, $\mathcal{F}$ returns a forgery $(M_1^*,(s_1^*,e_1^*))$ where $e_1^*=H(M_1^*\|r_1^*)$.


There is a non-negligible probability (at least $1/q$, but this has to be argued rigorously) that in Round 2 too $\mathcal{F}$ forges at the critical point (i.e., $M_1^*=M_0^*$ and $r_1^*=r_0^*$) but that the responses to the random oracle query $H(M_0^*\|r_0^*)$ were different (i.e., $e_1^*\neq e_0^*$). The intuitive reason is that $\mathcal{F}$ has to forge on some query, and in the worst case it chooses this point randomly. If this is indeed the case, then 
  $$s_0^*=k_0^*-\alpha e_0^* \text{ and } s_1^*=k_0^*-\alpha e_1^*,$$
(as $r_1^*=r_0^*$, but $e_1^*\neq e_0^*$) and $\mathcal{B}$ can solve for $\alpha$ 
 $$\alpha=\frac{s_1^*-s_0^*}{e_0^*-e_1^*}.$$

The above argument can be strengthened to accommodate a general $(\epsilon,q)$-adversary in the EU-CMA model. To simulate the signing oracle (for the EU-CMA model) the reduction only has to program the random oracle appropriately.$^3$. Bounding the success probability of the reduction for the  general adversary (that is successful with a non-negligible probability $\epsilon$) is quite technical and uses the so-called forking lemma. To be precise, it is shown in [PS00] that an $(\epsilon,q)$-forger $\mathcal{F}$ that breaks the Schnorr signature in the EU-CMA model can be used to break the DLP with probability $O(\epsilon^2/q)$.$^4$

Footnotes.

$^1$The DLP on a cyclic group $(G,g,q)$ requires finding $\alpha\in\mathbb{Z}_q$ given $g^\alpha\in G$.

$^2$That is, the queries up to the critical point are answered consistently, but the fresh queries after the critical point are answered independently of Round 1.

$^3$To generate a signature for a message $M$, select $e,k\in_R\mathbb{Z}_q$, set $r=(g^\alpha)^e\cdot g^k$, and program the random oracle to set $H(M\|r)=e$. Return $(M,(e,k))$ as the signature. It is not difficult to see that the message is valid, and from the right distribution.

$^4$It was later shown in a series of works [PV05,GBL08,Seu12] that the loss is tightness of $\epsilon/q$ is inherent (conditioned on the assumption that the so-called one-more discrete-logarithm is hard).

References.

[PS00] David Pointcheval and Jacques Stern. Security arguments for digital signatures and blind signatures. Journal of Cryptology, 2000.

[PV05] Pascal Paillier and Damien Vergnaud. Discrete-Log-Based Signatures May Not Be Equivalent to Discrete-Log. ASIACRYPT 2005.

[GBL08] Sanjam Garg, Raghav Bhaskar, and Satyanarayana V. Lokam. Improved Bounds on Security Reductions for Discrete-Log Based Signatures. CRYPTO 2008.

[Seu12] Yannich Seurin. On the Exact Security of Schnorr-Type Signatures in the Random Oracle Model. EUROCRYPT 2012.